Timeslot arbitration scheme

ABSTRACT

A method for arbitrating between a plurality of access requests issued in relation to a resource by a plurality of requestors, wherein each request can be one of at least two types, a first of the types having a higher latency associated with its performance than at least some of the other types, the method including the steps of: (a) receiving a plurality of the access requests; (the requests are not placed anywhere, they are simply received); (b) maintaining a current pointer that points to a current timeslot in a timeslot list, and at least one lookahead pointer that points to a future timeslot in the timeslot list; and (c) in the event an access request as arbitrated via the lookahead pointer is of the first type, initiating performance of the access request earlier than the position in the list suggests it would be performed should it be started when the current pointer reached the timeslot.

FIELD OF INVENTION

The present invention relates to a timeslot scheme for arbitrating between requests for access to a resource from multiple peripheral entities in a system.

The invention has primarily been developed for use in a printer controller chip that controls a printhead comprising one or more printhead modules constructed using microelectromechanical systems (MEMS) techniques, and will be described with reference to this application. However, it will be appreciated that the invention can be applied to other types of printing technologies in which analogous problems are faced.

BACKGROUND OF INVENTION

Manufacturing a printhead that has relatively high resolution and print-speed raises a number of problems.

Difficulties in manufacturing pagewidth printheads of any substantial size arise due to the relatively small dimensions of standard silicon wafers that are used in printhead (or printhead module) manufacture. For example, if it is desired to make an 8 inch wide pagewidth printhead, only one such printhead can be laid out on a standard 8-inch wafer, since such wafers are circular in plan. Manufacturing a pagewidth printhead from two or more smaller modules can reduce this limitation to some extent, but raises other problems related to providing a joint between adjacent printhead modules that is precise enough to avoid visible artefacts (which would typically take the form of noticeable lines) when the printhead is used. The problem is exacerbated in relatively high-resolution applications because of the tight tolerances dictated by the small spacing between nozzles.

The quality of a joint region between adjacent printhead modules relies on factors including a precision with which the abutting ends of each module can be manufactured, the accuracy with which they can be aligned when assembled into a single printhead, and other more practical factors such as management of ink channels behind the nozzles. It will be appreciated that the difficulties include relative vertical displacement of the printhead modules with respect to each other.

Whilst some of these issues may be dealt with by careful design and manufacture, the level of precision required renders it relatively expensive to manufacture printheads within the required tolerances. It would be desirable to provide a solution to one or more of the problems associated with precision manufacture and assembly of multiple printhead modules to form a printhead, and especially a pagewidth printhead.

In some cases, it is desirable to produce a number of different printhead module types or lengths on a substrate to maximise usage of the substrate's surface area. However, different sizes and types of modules will have different numbers and layouts of print nozzles, potentially including different horizontal and vertical offsets. Where two or more modules are to be joined to form a single printhead, there is also the problem of dealing with different seam shapes between abutting ends of joined modules, which again may incorporate vertical or horizontal offsets between the modules. Printhead controllers are usually dedicated application specific integrated circuits (ASICs) designed for specific use with a single type of printhead module, that is used by itself rather than with other modules. It would be desirable to provide a way in which different lengths and types of printhead modules could be accounted for using a single printer controller.

Printer controllers face other difficulties when two or more printhead modules are involved, especially if it is desired to send dot data to each of the printheads directly (rather than via a single printhead connected to the controller). One concern is that data delivered to different length controllers at the same rate will cause the shorter of the modules to be ready for printing before any longer modules. Where there is little difference involved, the issue may not be of importance, but for large length differences, the result is that the bandwidth of a shared memory from which the dot data is supplied to the modules is effectively left idle once one of the modules is full and the remaining module or modules is still being filled. It would be desirable to provide a way of improving memory bandwidth usage in a system comprising a plurality of printhead modules of uneven length.

In any printing system that includes multiple nozzles on a printhead or printhead module, there is the possibility of one or more of the nozzles failing in the field, or being inoperative due to manufacturing defect. Given the relatively large size of a typical printhead module, it would be desirable to provide some form of compensation for one or more “dead” nozzles. Where the printhead also outputs fixative on a per-nozzle basis, it is also desirable that the fixative is provided in such a way that dead nozzles are compensated for.

A printer controller can take the form of an integrated circuit, comprising a processor and one or more peripheral hardware units for implementing specific data manipulation functions. A number of these units and the processor may need access to a common resource such as memory. One way of arbitrating between multiple access requests for a common resource is timeslot arbitration, in which access to the resource is guaranteed to a particular requester during a predetermined timeslot.

One difficulty with this arrangement lies in the fact that not all access requests make the same demands on the resource in terms of timing and latency. For example, a memory read requires that data be fetched from memory, which may take a number of cycles, whereas a memory write can commence immediately. Timeslot arbitration does not take into account these differences, which may result in accesses being performed in a less efficient manner than might otherwise be the case. It would be desirable to provide a timeslot arbitration scheme that improved this efficiency as compared with prior art timeslot arbitration schemes.

Also of concern when allocating resources in a timeslot arbitration scheme is the fact that the priority of an access request may not be the same for all units. For example, it would be desirable to provide a timeslot arbitration scheme in which one requestor (typically the memory) is granted special priority such that its requests are dealt with earlier than would be the case in the absence of such priority.

In systems that use a memory and cache, a cache miss (in which an attempt to load data or an instruction from a cache fails) results in a memory access followed by a cache update. It is often desirable when updating the cache in this way to update data other than that which was actually missed. A typical example would be a cache miss for a byte resulting in an entire word or line of the cache associated with that byte being updated. However, this can have the effect of tying up bandwidth between the memory (or a memory manager) and the processor where the bandwidth is such that several cycles are required to transfer the entire word or line to the cache. It would be desirable to provide a mechanism for updating a cache that improved cache update speed and/or efficiency.

Most integrated circuits an externally provided signal as (or to generate) a clock, often provided from a dedicated clock generation circuit. This is often due to the difficulties of providing an onboard clock that can operate at a speed that is predictable. Manufacturing tolerances of such on-board clock generation circuitry can result in clock rates that vary by a factor of two, and operating temperatures can increase this margin by an additional factor of two. In some cases, the particular rate at which the clock operates is not of particular concern. However, where the integrated circuit will be writing to an internal circuit that is sensitive to the time over which a signal is provided, it may be undesirable to have the signal be applied for too long or short a time. For example, flash memory is sensitive to being written too for too long a period. It would be desirable to provide a mechanism for adjusting a rate of an on-chip system clock to take into account the impact of manufacturing variations on clockspeed.

One form of attacking a secure chip is to induce (usually by increasing) a clock speed that takes the logic outside its rated operating frequency. One way of doing this is to reduce the temperature of the integrated circuit, which can cause the clock to race. Above a certain frequency, some logic will start malfunctioning. In some cases, the malfunction can be such that information on the chip that would otherwise be secure may become available to an external connection. It would be desirable to protect an integrated circuit from such attacks.

In an integrated circuit comprising non-volatile memory, a power failure can result in unintentional behaviour. For example, if an address or data becomes unreliable due to falling voltage supplied to the circuit but there is still sufficient power to cause a write, incorrect data can be written. Even worse, the data (incorrect or not) could be written to the wrong memory. The problem is exacerbated with multi-word writes. It would be desirable to provide a mechanism for reducing or preventing spurious writes when power to an integrated circuit is failing.

In an integrated circuit, it is often desirable to reduce unauthorised access to the contents of memory. This is particularly the case where the memory includes a key or some other form of security information that allows the integrated circuit to communicate with another entity (such as another integrated circuit, for example) in a secure manner. It would be particularly advantageous to prevent attacks involving direct probing of memory addresses by physically investigating the chip (as distinct from electronic or logical attacks via manipulation of signals and power supplied to the integrated circuit).

It is also desirable to provide an environment where the manufacturer of the integrated circuit (or some other authorised entity) can verify or authorize code to be run on an integrated circuit.

Another desideratum would be the ability of two or more entities, such as integrated circuits, to communicate with each other in a secure manner. It would also be desirable to provide a mechanism for secure communication between a first entity and a second entity, where the two entities, whilst capable of some form of secure communication, are not able to establish such communication between themselves.

In a system that uses resources (such as a printer, which uses inks) it may be desirable to monitor and update a record related to resource usage. Authenticating ink quality can be a major issue, since the attributes of inks used by a given printhead can be quite specific. Use of incorrect ink can result in anything from misfiring or poor performance to damage or destruction of the printhead. It would therefore be desirable to provide a system that enables authentication of the correct ink being used, as well as providing various support systems secure enabling refilling of ink cartridges.

In a system that prevents unauthorized programs from being loaded onto or run on an integrated circuit, it can be laborious to allow developers of software to access the circuits during software development. Enabling access to integrated circuits of a particular type requires authenticating software with a relatively high-level key. Distributing the key for use by developers is inherently unsafe, since a single leak of the key outside the organization could endanger security of all chips that use a related key to authorize programs. Having a small number of people with high-security clearance available to authenticate programs for testing can be inconvenient, particularly in the case where frequent incremental changes in programs during development require testing. It would be desirable to provide a mechanism for allowing access to one or more integrated circuits without risking the security of other integrated circuits in a series of such integrated circuits.

In symmetric key security, a message, denoted by M, is plaintext. The process of transforming M into ciphertext C, where the substance of M is hidden, is called encryption. The process of transforming C back into M is called decryption. Referring to the encryption function as E, and the decryption function as D, we have the following identities: E[M]=C D[C]=M

Therefore the following identity is true: D[E[M]]=M

A symmetric encryption algorithm is one where:

-   -   the encryption function E relies on key K₁,     -   the decryption function D relies on key K₂,     -   K₂ can be derived from K₁, and     -   K₁ can be derived from K₂,

In most symmetric algorithms, K₁ equals K₂. However, even if K₁ does not equal K₂, given that one key can be derived from the other, a single key K can suffice for the mathematical definition. Thus: E_(K)[M]=C D_(K)[C]=M

The security of these algorithms rests very much in the key K. Knowledge of K allows anyone to encrypt or decrypt. Consequently K must remain a secret for the duration of the value of M. For example, M may be a wartime message “My current position is grid position 123-456”. Once the war is over the value of M is greatly reduced, and if K is made public, the knowledge of the combat unit's position may be of no relevance whatsoever. The security of the particular symmetric algorithm is a function of two things: the strength of the algorithm and the length of the key.

An asymmetric encryption algorithm is one where:

-   -   the encryption function E relies on key K₁,     -   the decryption function D relies on key K₂,     -   K₂ cannot be derived from K₁ in a reasonable amount of time, and     -   K₁ cannot be derived from K₂ in a reasonable amount of time.

Thus: E_(K1)[M]=C D_(K2)[C]=M

These algorithms are also called public-key because one key K₁ can be made public. Thus anyone can encrypt a message (using K₁) but only the person with the corresponding decryption key (K₂) can decrypt and thus read the message.

In most cases, the following identity also holds: E_(K2)[M]=C D_(K1)[C]=M

This identity is very important because it implies that anyone with the public key K₁ can see M and know that it came from the owner of K₂. No-one else could have generated C because to do so would imply knowledge of K₂. This gives rise to a different application, unrelated to encryption—digital signatures.

A number of public key cryptographic algorithms exist. Most are impractical to implement, and many generate a very large C for a given M or require enormous keys. Still others, while secure, are far too slow to be practical for several years. Because of this, many public key systems are hybrid—a public key mechanism is used to transmit a symmetric session key, and then the session key is used for the actual messages.

All of the algorithms have a problem in terms of key selection. A random number is simply not secure enough. The two large primes p and q must be chosen carefully—there are certain weak combinations that can be factored more easily (some of the weak keys can be tested for). But nonetheless, key selection is not a simple matter of randomly selecting 1024 bits for example. Consequently the key selection process must also be secure.

Symmetric and asymmetric schemes both suffer from a difficulty in allowing establishment of multiple relationships between one entity and a two or more others, without the need to provide multiple sets of keys. For example, if a main entity wants to establish secure communications with two or more additional entities, it will need to maintain a different key for each of the additional entities. For practical reasons, it is desirable to avoid generating and storing large numbers of keys. To reduce key numbers, two or more of the entities may use the same key to communicate with the main entity. However, this means that the main entity cannot be sure which of the entities it is communicating with. Similarly, messages from the main entity to one of the entities can be decrypted by any of the other entities with the same key. It would be desirable if a mechanism could be provided to allow secure communication between a main entity and one or more other entities that overcomes at least some of the shortcomings of prior art.

In a system where a first entity is capable of secure communication of some form, it may be desirable to establish a relationship with another entity without providing the other entity with any information related the first entity's security features. Typically, the security features might include a key or a cryptographic function. It would be desirable to provide a mechanism for enabling secure communications between a first and second entity when they do not share the requisite secret function, key or other relationship to enable them to establish trust.

A number of other aspects, features, preferences and embodiments are disclosed in the Detailed Description of the Preferred Embodiment below.

SUMMARY OF THE INVENTION

In accordance with a first aspect of the invention, there is provided a method for arbitrating between a plurality of access requests issued in relation to a resource by a plurality of requesters, wherein each request can be one of at least two types, a first of the types having a higher latency associated with its performance than at least some of the other types, the method including the steps of:

-   (a) receiving a plurality of the access requests; (the requests are     not placed anywhere, they are simply received) -   (b) maintaining a current pointer that points to a current timeslot     in a timeslot list, and at least one lookahead pointer that points     to a future timeslot in the timeslot list; and -   (c) in the event an access request as arbitrated via the lookahead     pointer is of the first type, initiating performance of the access     request earlier than the position in the list suggests it would be     performed should it be started when the current pointer reached the     timeslot.

Preferably, step (c) includes the substep of performing the access request indicated by the lookahead pointer immediately after the access request indicated by the current pointer is performed.

It is preferred that step (c) includes the substep of performing the access request indicated by the lookahead pointer immediately after the access request indicated by the current pointer is performed.

In a preferred embodiment, the number of timeslots between the timeslot indicated by the lookahead pointer and the timeslot indicated by the current pointer takes into account a latency difference between performing an access request of the first type and at least one of the other access request types.

In accordance with a second aspect of the invention, there is provided an integrated circuit including:

-   -   a plurality of operative units, each of which is capable of         issuing a request for access to a memory accessible by the         integrated circuit; and     -   an timeslot arbitrator for arbitrating between requests issued         by the operative units for access to the memory, wherein each         request can be one of at least two types, a first of the types         having a higher latency associated with its performance than at         least some of the other types, the timeslot arbitrator being         configured to:

-   (a) receive a plurality of the access requests;

-   (b) maintain a current pointer that points to a current timeslot in     a timeslot list, and at least one lookahead pointer that points to a     future timeslot in the timeslot list; and

-   (c) in the event the access request as arbitrated via the lookahead     pointer is of the first type, performing the access request earlier     than the position in the list suggests it should be performed should     it be started when the current pointer reached the timeslot.

Preferably, the first type of access request is a memory write request.

Preferably, the integrated circuit includes a memory interface unit operatively connected with, and under the control of, the timeslot arbitrator, and wherein the memory interface is operatively connected to:

-   -   one or more of the operative units via one or more         communications buses, and     -   the memory via a memory bus of greater width than the         communications buses.

In a preferred form, the number of timeslots between the timeslot indicated by the lookahead pointer and the timeslot indicated by the current pointer takes into account a latency difference between performing an access request of the first type and at least one of access request types.

In accordance with a third aspect of the invention, there is provided a method of arbitrating between access requests from a plurality of requestors for access to a resource, wherein at least one of the requesters is defined as higher priority access to the resource, the method comprising the steps of:

-   (a) receiving a plurality of the access requests; -   (b) in the event an access request from the at least one of the     requestors is received, initiating performance of the access request     in preference to the requestor as specified by the timeslot list and     regardless of whether or not the at least one of the requestors is     in the timeslot list.

Preferably, the at least one requestor requires lower latency access to the resource than at least one of the other requestors from which access requests can be received.

Preferably, the at least one requestor is a processor and/or the resource is a memory.

Preferably, step (b) includes the substep of performing the access request from the requester immediately following completion of any current access request being reformed.

In a preferred form, step (b) is performed such that a frequency of the at least one requestor being granted preferential performance of its access requests is limited within a time period. More preferably, early performance of access requests from the at least one requestor is restricted to a maximum number of times within a predetermined number of timeslots.

Preferably, the requesters are hardware units on an integrated circuit and the method is implemented by a timeslot arbitrator unit on the integrated circuit.

In accordance with a fourth aspect of the invention, there is provided a method according to the third aspect, wherein each request can also be one of at least two types, a first of the types having a higher latency associated with its performance than at least some of the other types, the method including the steps of:

-   (a) receiving a plurality of the access requests; -   (b) maintaining a current pointer that points to a current timeslot     in the timeslot list, and at least one lookahead pointer that points     to a future timeslot in the timeslot list; and -   (c) in the event an access request as arbitrated via the lookahead     pointer is of the first type, initiating performance of the access     request earlier than its position in the list suggests it should be     performed should it be started when the current pointer reached the     timeslot.

BRIEF DESCRIPTION OF THE DRAWINGS

Preferred and other embodiments of the invention will now be described, by way of example only, with reference to the accompanying drawings, in which:

FIG. 1 is an example of state machine notation

FIG. 2 shows document data flow in a printer

FIG. 3 is an example of a single printer controller (hereinafter “SoPEC”) A4 simplex printer system

FIG. 4 is an example of a dual SoPEC A4 duplex printer system

FIG. 5 is an example of a dual SoPEC A3 simplex printer system

FIG. 6 is an example of a quad SoPEC A3 duplex printer system

FIG. 7 is an example of a SoPEC A4 simplex printing system with an extra SoPEC used as DRAM storage

FIG. 8 is an example of an A3 duplex printing system featuring four printing SoPECs

FIG. 9 shows pages containing different numbers of bands

FIG. 10 shows the contents of a page band

FIG. 11 illustrates a page data path from host to SoPEC

FIG. 12 shows a page structure

FIG. 13 shows a SoPEC system top level partition

FIG. 14 shows a SoPEC CPU memory map (not to scale)

FIG. 15 is a block diagram of CPU

FIG. 16 shows CPU bus transactions

FIG. 17 shows a state machine for a CPU subsystem slave

FIG. 18 shows a SoPEC CPU memory map (not to scale)

FIG. 19 shows an external signal view of a memory management unit (hereinafter “MMU”) sub-block partition

FIG. 20 shows an internal signal view of an MMU sub-block partition

FIG. 21 shows a DRAM write buffer

FIG. 22 shows DIU waveforms for multiple transactions

FIG. 23 shows a SoPEC LEON CPU core

FIG. 24 shows a cache data RAM wrapper

FIG. 25 shows a realtime debug unit block diagram

FIG. 26 shows interrupt acknowledge cycles for single and pending interrupts

FIG. 27 shows an A3 duplex system featuring four printing SoPECs with a single SoPEC DRAM device

FIG. 28 is an SCB block diagram

FIG. 29 is a logical view of the SCB of FIG. 28

FIG. 30 shows an ISI configuration with four SoPEC devices

FIG. 31 shows half-duplex interleaved transmission from ISIMaster to ISISlave

FIG. 32 shows ISI transactions

FIG. 33 shows an ISI long packet

FIG. 34 shows an ISI ping packet

FIG. 35 shows a short ISI packet

FIG. 36 shows successful transmission of two long packets with sequence bit toggling

FIG. 37 shows sequence bit operation with errored long packet

FIG. 38 shows sequence bit operation with ACK error

FIG. 39 shows an ISI sub-block partition

FIG. 40 shows an ISI serial interface engine functional block diagram

FIG. 41 is an SIE edge detection and data 10 diagram

FIG. 42 is an SIE Rx/Tx state machine Tx cycle state diagram

FIG. 43 shows an SIE Rx/Tx state machine Tx bit stuff ‘0’ cycle state diagram

FIG. 44 shows an SIE Rx/Tx state machine Tx bit stuff ‘1’ cycle state diagram

FIG. 45 shows an SIE Rx/Tx state machine Rx cycle state diagram

FIG. 46 shows an SIE Tx functional timing example

FIG. 47 shows an SIE Rx functional timing example

FIG. 48 shows an SIE Rx/Tx FIFO block diagram

FIG. 49 shows SIE Rx/Tx FIFO control signal gating

FIG. 50 shows an SIE bit stuffing state machine Tx cycle state diagram

FIG. 51 shows an SIE bit stripping state machine Rx cycle state diagram

FIG. 52 shows a CRC16 generation/checking shift register

FIG. 53 shows circular buffer operation

FIG. 54 shows duty cycle select

FIG. 55 shows a GPIO partition

FIG. 56 shows a motor control RTL diagram

FIG. 57 is an input de-glitch RTL diagram

FIG. 58 is a frequency analyser RTL diagram

FIG. 59 shows a brushless DC controller

FIG. 60 shows a period measure unit

FIG. 61 shows line synch generation logic

FIG. 62 shows an ICU partition

FIG. 63 is an interrupt clear state diagram

FIG. 63A Timers sub-block partition diagram

FIG. 64 is a watchdog timer RTL diagram

FIG. 65 is a generic timer RTL diagram

FIG. 66 is a schematic of a timing pulse generator

FIG. 67 is a Pulse generator RTL diagram

FIG. 68 shows a SoPEC clock relationship

FIG. 69 shows a CPR block partition

FIG. 70 shows reset deglitch logic

FIG. 71 shows reset synchronizer logic

FIG. 72 is a clock gate logic diagram

FIG. 73 shows a PLL and Clock divider logic

FIG. 74 shows a PLL control state machine diagram

FIG. 75 shows a LSS master system-level interface

FIG. 76 shows START and STOP conditions

FIG. 77 shows an LSS transfer of 2 data bytes

FIG. 78 is an example of an LSS write to a QA Chip

FIG. 79 is an example of an LSS read from QA Chip

FIG. 80 shows an LSS block diagram

FIG. 81 shows an LSS multi-command transaction

FIG. 82 shows start and stop generation based on previous bus state

FIG. 83 shows an LSS master state machine

FIG. 84 shows LSS master timing

FIG. 85 shows a SoPEC system top level partition

FIG. 86 shows an ead bus with 3 cycle random DRAM read accesses

FIG. 87 shows interleaving of CPU and non-CPU read accesses

FIG. 88 shows interleaving of read and write accesses with 3 cycle random DRAM accesses

FIG. 89 shows interleaving of write accesses with 3 cycle random DRAM accesses

FIG. 90 shows a read protocol for a SoPEC Unit making a single 256-bit access

FIG. 91 shows a read protocol for a SoPEC Unit making a single 256-bit access

FIG. 92 shows a write protocol for a SoPEC Unit making a single 256-bit access

FIG. 93 shows a protocol for a posted, masked, 128-bit write by the CPU

FIG. 94 shows a write protocol shown for CDU making four contiguous 64-bit accesses

FIG. 95 shows timeslot-based arbitration

FIG. 96 shows timeslot-based arbitration with separate pointers

FIG. 97 shows a first example (a) of separate read and write arbitration

FIG. 98 shows a second example (b) of separate read and write arbitration

FIG. 99 shows a third example (c) of separate read and write arbitration

FIG. 100 shows a DIU partition

FIG. 101 shows a DIU partition

FIG. 102 shows multiplexing and address translation logic for two memory instances

FIG. 103 shows a timing of dau_dcu_valID, dcu_dau_adv and dcu_dau_wadv

FIG. 104 shows a DCU state machine

FIG. 105 shows random read timing

FIG. 106 shows random write timing

FIG. 107 shows refresh timing

FIG. 108 shows page mode write timing

FIG. 109 shows timing of non-CPU DIU read access

FIG. 110 shows timing of CPU DIU read access

FIG. 111 shows a CPU DIU read access

FIG. 112 shows timing of CPU DIU write access

FIG. 113 shows timing of a non-CDU/non-CPU DIU write access

FIG. 114 shows timing of CDU DIU write access

FIG. 115 shows command multiplexor sub-block partition

FIG. 116 shows command multiplexor timing at DIU requesters interface

FIG. 117 shows generation of re_arbitrate and re_arbitrate_wadv

FIG. 118 shows CPU interface and arbitration logic

FIG. 119 shows arbitration timing

FIG. 120 shows setting RotationSync to enable a new rotation.

FIG. 121 shows a timeslot based arbitration

FIG. 122 shows a timeslot based arbitration with separate pointers

FIG. 123 shows a CPU pre-access write lookahead pointer

FIG. 124 shows arbitration hierarchy

FIG. 125 shows hierarchical round-robin priority comparison

FIG. 126 shows a read multiplexor partition

FIG. 127 shows a read command queue (4 deep buffer)

FIG. 128 shows state-machines for shared read bus accesses

FIG. 129 shows a write multiplexor partition

FIG. 130 shows a read multiplexer timing for back-to-back shared read bus transfer

FIG. 131 shows a write multiplexer partition

FIG. 132 shows a block diagram of a PCU

FIG. 133 shows PCU accesses to PEP registers

FIG. 134 shows command arbitration and execution

FIG. 135 shows DRAM command access state machine

FIG. 136 shows an outline of contone data flow with respect to CDU

FIG. 137 shows a DRAM storage arrangement for a single line of JPEG 8×8 blocks in 4 colors

FIG. 138 shows a read control unit state machine

FIG. 139 shows a memory arrangement of JPEG blocks

FIG. 140 shows a contone data write state machine

FIG. 141 shows lead-in and lead-out clipping of contone data in multi-SoPEC environment

FIG. 142 shows a block diagram of CFU

FIG. 143 shows a DRAM storage arrangement for a single line of JPEG blocks in 4 colors

FIG. 144 shows a block diagram of color space converter

FIG. 145 shows a converter/invertor

FIG. 146 shows a high-level block diagram of LBD in context

FIG. 147 shows a schematic outline of the LBD and the SFU

FIG. 148 shows a block diagram of lossless bi-level decoder

FIG. 149 shows a stream decoder block diagram

FIG. 150 shows a command controller block diagram

FIG. 151 shows a state diagram for command controller (CC) state machine

FIG. 152 shows a next edge unit block diagram

FIG. 153 shows a next edge unit buffer diagram

FIG. 154 shows a next edge unit edge detect diagram

FIG. 155 shows a state diagram for the next edge unit state machine

FIG. 156 shows a line fill unit block diagram

FIG. 157 shows a state diagram for the Line Fill Unit (LFU) state machine

FIG. 158 shows a bi-level DRAM buffer

FIG. 159 shows interfaces between LBD/SFU/HCU

FIG. 160 shows an SFU sub-block partition

FIG. 161 shows an LBDPrevLineFifo sub-block

FIG. 162 shows timing of signals on the LBDPrevLineFIFO interface to DIU and address generator

FIG. 163 shows timing of signals on LBDPrevLineFIFO interface to DIU and address generator

FIG. 164 shows LBDNextLineFifo sub-block

FIG. 165 shows timing of signals on LBDNextLineFIFO interface to DIU and address generator

FIG. 166 shows LBDNextLineFIFO DIU interface state diagram

FIG. 167 shows an LDB to SFU write interface

FIG. 168 shows an LDB to SFU read interface (within a line)

FIG. 169 shows an HCUReadLineFifo Sub-block

FIG. 170 shows a DIU write Interface

FIG. 171 shows a DIU Read Interface multiplexing by select_hrfplf

FIG. 172 shows DIU read request arbitration logic

FIG. 173 shows address generation

FIG. 174 shows an X scaling control unit

FIG. 175 Y shows a scaling control unit

FIG. 176 shows an overview of X and Y scaling at HCU interface

FIG. 177 shows a high level block diagram of TE in context

FIG. 178 shows a QR Code

FIG. 179 shows Netpage tag structure

FIG. 180 shows a Netpage tag with data rendered at 1600 dpi (magnified view)

FIG. 181 shows an example of 2×2 dots for each block of QR code

FIG. 182 shows placement of tags for portrait & landscape printing

FIG. 183 shows agGeneral representation of tag placement

FIG. 184 shows composition of SoPEC's tag format structure

FIG. 185 shows a simple 3×3 tag structure

FIG. 186 shows 3×3 tag redesigned for 21×21 area (not simple replication)

FIG. 187 shows a TE Block Diagram

FIG. 188 shows a TE Hierarchy

FIG. 189 shows a block diagram of PCU accesses

FIG. 190 shows a tag encoder top-level FSM

FIG. 191 shows generated control signals

FIG. 192 shows logic to combine dot information and encoded data

FIG. 193 shows generation of Lastdotintag/1

FIG. 194 shows generation of Dot Position Valid

FIG. 195 shows generation of write enable to the TFU

FIG. 196 shows generation of Tag Dot Number

FIG. 197 shows TDI Architecture

FIG. 198 shows data flow through the TDI

FIG. 199 shows raw tag data interface block diagram

FIG. 200 shows an RTDI State Flow Diagram

FIG. 201 shows a relationship between TE_endoftagdata, cdu_startofbandstore and cdu_endofbandstore

FIG. 202 shows a TDi State Flow Diagram

FIG. 203 shows mapping of the tag data to codewords 0-7

FIG. 204 shows coding and mapping of uncoded fixed tag data for (15,5) RS encoder

FIG. 205 shows mapping of pre-coded fixed tag data

FIG. 206 shows coding and mapping of variable tag data for (15,7) RS encoder

FIG. 207 shows coding and mapping of uncoded fixed tag data for (15,7) RS encoder

FIG. 208 shows mapping of 2D decoded variable tag data

FIG. 209 shows a simple block diagram for an m=4 Reed Solomon encoder

FIG. 210 shows an RS encoder I/O diagram

FIG. 211 shows a (15,5) & (15,7) RS encoder block diagram

FIG. 212 shows a (15,5) RS encoder timing diagram

FIG. 213 shows a (15,7) RS encoder timing diagram

FIG. 214 shows a circuit for multiplying by alpha³

FIG. 215 shows adding two field elements

FIG. 216 shows an RS encoder implementation

FIG. 217 shows an encoded tag data interface

FIG. 218 shows an encoded fixed tag data interface

FIG. 219 shows an encoded variable tag data interface

FIG. 220 shows an encoded variable tag data sub-buffer

FIG. 221 shows a breakdown of the tag format structure

FIG. 222 shows a TFSI FSM state flow diagram

FIG. 223 shows a TFS block diagram

FIG. 224 shows a table A interface block diagram

FIG. 225 shows a table A address generator

FIG. 226 shows a table C interface block diagram

FIG. 227 shows a table B interface block diagram

FIG. 228 shows interfaces between TE, TFU and HCU

FIG. 229 shows a 16-byte FIFO in TFU

FIG. 230 shows a high level block diagram showing the HCU and its external interfaces

FIG. 231 shows a block diagram of the HCU

FIG. 232 shows a block diagram of the control unit

FIG. 233 shows a block diagram of determine advdot unit

FIG. 234 shows a page structure

FIG. 235 shows a block diagram of a margin unit

FIG. 236 shows a block diagram of a dither matrix table interface

FIG. 237 shows an example of reading lines of dither matrix from DRAM

FIG. 238 shows a state machine to read dither matrix table

FIG. 239 shows a contone dotgen unit

FIG. 240 shows a block diagram of dot reorg unit

FIG. 241 shows an HCU to DNC interface (also used in DNC to DWU, LLU to PHI)

FIG. 242 shows SFU to HCU interface (all feeders to HCU)

FIG. 243 shows representative logic of the SFU to HCU interface

FIG. 244 shows a high-level block diagram of DNC

FIG. 245 shows a dead nozzle table format

FIG. 246 shows set of dots operated on for error diffusion

FIG. 247 shows a block diagram of DNC

FIG. 248 shows a sub-block diagram of ink replacement unit

FIG. 249 shows a dead nozzle table state machine

FIG. 250 shows logic for dead nozzle removal and ink replacement

FIG. 251 shows a sub-block diagram of error diffusion unit

FIG. 252 shows a maximum length 32-bit LFSR used for random bit generation

FIG. 253 shows a high-level data flow diagram of DWU in context

FIG. 254 shows a printhead nozzle layout for 36-nozzle bi-lithic printhead

FIG. 255 shows a printhead nozzle layout for a 36-nozzle bi-lithic printhead

FIG. 256 shows a dot line store logical representation

FIG. 257 shows a conceptual view of printhead row alignment

FIG. 258 shows a conceptual view of printhead rows (as seen by the LLU and PHI)

FIG. 259 shows a comparison of 1.5×v 2× buffering

FIG. 260 shows an even dot order in DRAM (increasing sense, 13320 dot wide line)

FIG. 261 shows an even dot order in DRAM (decreasing sense, 13320 dot wide line)

FIG. 262 shows a dotline FIFO data structure in DRAM

FIG. 263 shows a DWU partition

FIG. 264 shows a buffer address generator sub-block

FIG. 265 shows a DIU Interface sub-block

FIG. 266 shows an interface controller state diagram

FIG. 267 shows a high level data flow diagram of LLU in context

FIG. 268 shows paper and printhead nozzles relationship (example with D₁=D₂=5)

FIG. 269 shows printhead structure and dot generate order

FIG. 270 shows an order of dot data generation and transmission

FIG. 271 shows a conceptual view of printhead rows

FIG. 272 shows a dotline FIFO data structure in DRAM (LLU specification)

FIG. 273 shows an LLU partition

FIG. 274 shows a dot generator RTL diagram

FIG. 275 shows a DIU interface

FIG. 276 shows an interface controller state diagram

FIG. 277 shows high-level data flow diagram of PHI in context

FIG. 278 shows power on reset

FIG. 279 shows printhead data rate equalization

FIG. 280 shows a printhead structure and dot generate order

FIG. 281 shows an order of dot data generation and transmission

FIG. 282 shows an order of dot data generation and transmission (single printhead case)

FIG. 283 shows printhead interface timing parameters

FIG. 284 shows printhead timing with margining

FIG. 285 shows a PHI block partition

FIG. 286 shows a sync generator state diagram

FIG. 287 shows a line sync de-glitch RTL diagram

FIG. 288 shows a fire generator state diagram

FIG. 289 shows a PHI controller state machine

FIG. 290 shows a datapath unit partition

FIG. 291 shows a dot order controller state diagram

FIG. 292 shows a data generator state diagram

FIG. 293 shows data serializer timing

FIG. 294 shows a data serializer RTL Diagram

FIG. 295 shows printhead types 0 to 7

FIG. 296 shows an ideal join between two dilithic printhead segments

FIG. 297 shows an example of a join between two bilithic printhead segments

FIG. 298 shows printable vs non-printable area under new definition (looking at colors as if 1 row only)

FIG. 299 shows identification of printhead nozzles and shift-register sequences for printheads in arrangement 1

FIG. 300 shows demultiplexing of data within the printheads in arrangement 1

FIG. 301 shows double data rate signalling for a type 0 printhead in arrangement 1

FIG. 302 shows double data rate signalling for a type 1 printhead in arrangement 1

FIG. 303 shows identification of printheads nozzles and shift-register sequences for printheads in arrangement 2

FIG. 304 shows demultiplexing of data within the printheads in arrangement 2

FIG. 305 shows double data rate signalling for a type 0 printhead in arrangement 2

FIG. 306 shows double data rate signalling for a type 1 printhead in arrangement 2

FIG. 307 shows all 8 printhead arrangements

FIG. 308 shows a printhead structure

FIG. 309 shows a column Structure

FIG. 310 shows a printhead dot shift register dot mapping to page

FIG. 311 shows data timing during printing

FIG. 312 shows print quality

FIG. 313 shows fire and select shift register setup for printing

FIG. 314 shows a fire pattern across butt end of printhead chips

FIG. 315 shows fire pattern generation

FIG. 316 shows determination of select shift register value

FIG. 317 shows timing for printing signals

FIG. 318 shows initialisation of printheads

FIG. 319 shows a nozzle test latching circuit

FIG. 320 shows nozzle testing

FIG. 321 shows a temperature reading

FIG. 322 shows CMOS testing

FIG. 323 shows a reticle layout

FIG. 324 shows a stepper pattern on Wafer

FIG. 325 shows relationship between datasets

FIG. 326 shows a validation hierarchy

FIG. 327 shows development of operating system code

FIG. 328 shows protocol for directly verifying reads from ChipR

FIG. 329 shows a protocol for signature translation protocol

FIG. 330 shows a protocol for a direct authenticated write

FIG. 331 shows an alternative protocol for a direct authenticated write

FIG. 332 shows a protocol for basic update of permissions

FIG. 333 shows a protocol for a multiple-key update

FIG. 334 shows a protocol for a single-key authenticated read

FIG. 335 shows a protocol for a single-key authenticated write

FIG. 336 shows a protocol for a single-key update of permissions

FIG. 337 shows a protocol for a single-key update

FIG. 338 shows a protocol for a multiple-key single-M authenticated read

FIG. 339 shows a protocol for a multiple-key authenticated write

FIG. 340 shows a protocol for a multiple-key update of permissions

FIG. 341 shows a protocol for a multiple-key update

FIG. 342 shows a protocol for a multiple-key multiple-M authenticated read

FIG. 343 shows a protocol for a multiple-key authenticated write

FIG. 344 shows a protocol for a multiple-key update of permissions

FIG. 345 shows a protocol for a multiple-key update

FIG. 346 shows relationship of permissions bits to M[n] access bits

FIG. 347 shows 160-bit maximal period LFSR

FIG. 348 shows clock filter

FIG. 349 shows tamper detection line

FIG. 350 shows an oversize nMOS transistor layout of Tamper Detection Line

FIG. 351 shows a Tamper Detection Line

FIG. 352 shows how Tamper Detection Lines cover the Noise Generator

FIG. 353 shows a prior art FET Implementation of CMOS inverter

FIG. 354 shows non-flashing CMOS

FIG. 355 shows components of a printer-based refill device

FIG. 356 shows refilling of printers by printer-based refill device

FIG. 357 shows components of a home refill station

FIG. 358 shows a three-ink reservoir unit

FIG. 359 shows refill of ink cartridges in a home refill station

FIG. 360 shows components of a commercial refill station

FIG. 361 shows an ink reservoir unit

FIG. 362 shows refill of ink cartridges in a commercial refill station (showing a single refill unit)

FIG. 363 shows equivalent signature generation

FIG. 364 shows a basic field definition

FIG. 365 shows an example of defining field sizes and positions

FIG. 366 shows permissions

FIG. 367 shows a first example of permissions for a field

FIG. 368 shows a second example of permissions for a field

FIG. 369 shows field attributes

FIG. 370 shows an output signature generation data format for Read

FIG. 371 shows an input signature verification data format for Test

FIG. 372 shows an output signature generation data format for Translate

FIG. 373 shows an input signature verification data format for WriteAuth

FIG. 374 shows input signature data format for ReplaceKey

FIG. 375 shows a key replacement map

FIG. 376 shows a key replacement map after K₁ is replaced

FIG. 377 shows a key replacement process

FIG. 378 shows an output signature data format for GetProgramKey

FIG. 379 shows transfer and rollback process

FIG. 380 shows an upgrade flow

FIG. 381 shows authorised ink refill paths in the printing system

FIG. 382 shows an input signature verification data format for XferAmount

FIG. 383 shows a transfer and rollback process

FIG. 384 shows an upgrade flow

FIG. 385 shows authorised upgrade paths in the printing system

FIG. 386 shows a direct signature validation sequence

FIG. 387 shows signature validation using translation

FIG. 388 shows setup of preauth field attributes

FIG. 389 shows a high level block diagram of QA Chip

FIG. 390 shows an analogue unit

FIG. 391 shows a serial bus protocol for trimming

FIG. 392 shows a block diagram of a trim unit

FIG. 393 shows a block diagram of a CPU of the QA chip

FIG. 394 shows block diagram of an MIU

FIG. 395 shows a block diagram of memory components

FIG. 396 shows a first byte sent to an IOU

FIG. 397 shows a block diagram of the IOU

FIG. 398 shows a relationship between external SDa and SClk and generation of internal signals

FIG. 399 shows block diagram of ALU

FIG. 400 shows a block diagram of DataSel

FIG. 401 shows a block diagram of ROR

FIG. 402 shows a block diagram of the ALU's IO block

FIG. 403 shows a block diagram of PCU

FIG. 404 shows a block diagram of an Address Generator Unit

FIG. 405 shows a block diagram for a Counter Unit

FIG. 406 shows a block diagram of PMU

FIG. 407 shows a state machine for PMU

FIG. 408 shows a block diagram of MRU

FIG. 409 shows simplified MAU state machine

FIG. 410 shows power-on reset behaviour

FIG. 411 shows a ring oscillator block diagram

FIG. 412 shows a system clock duty cycle

DETAILED DESCRIPTION OF PREFERRED AND OTHER EMBODIMENTS

It will be appreciated that the detailed description that follows takes the form of a highly detailed design of the invention, including supporting hardware and software. A high level of detailed disclosure is provided to ensure that one skilled in the art will have ample guidance for implementing the invention.

Imperative phrases such as “must”, “requires”, “necessary” and “important” (and similar language) should be read as being indicative of being necessary only for the preferred embodiment actually being described. As such, unless the opposite is clear from the context, imperative wording should not be interpreted as such. Nothing in the detailed description is to be understood as limiting the scope of the invention, which is intended to be defined as widely as is defined in the accompanying claims.

Indications of expected rates, frequencies, costs, and other quantitative values are exemplary and estimated only, and are made in good faith. Nothing in this specification should be read as implying that a particular commercial embodiment is or will be capable of a particular performance level in any measurable area.

It will be appreciated that the principles, methods and hardware described throughout this document can be applied to other fields. Much of the security-related disclosure, for example, can be applied to many other fields that require secure communications between entities, and certainly has application far beyond the field of printers.

System Overview

The preferred of the present invention is implemented in a printer using microelectromechanical systems (MEMS) printheads. The printer can receive data from, for example, a personal computer such as an IBM compatible PC or Apple computer. In other embodiments, the printer can receive data directly from, for example, a digital still or video camera. The particular choice of communication link is not important, and can be based, for example, on USB, Firewire, Bluetooth or any other wireless or hardwired communications protocol.

Print System Overview

3 Introduction

This document describes the SoPEC (Small office home office Print Engine Controller) ASIC (Application Specific Integrated Circuit) suitable for use in, for example, SoHo printer products. The SoPEC ASIC is intended to be a low cost solution for bi-lithic printhead control, replacing the multichip solutions in larger more professional systems with a single chip. The increased cost competitiveness is achieved by integrating several systems such as a modified PEC1 printing pipeline, CPU control system, peripherals and memory sub-system onto one SoC ASIC, reducing component count and simplifying board design.

This section will give a general introduction to Memjet printing systems, introduce the components that make a bi-lithic printhead system, describe possible system architectures and show how several SoPECs can be used to achieve A3 and A4 duplex printing. The section “SoPEC ASIC” describes the SoC SoPEC ASIC, with subsections describing the CPU, DRAM and Print Engine Pipeline subsystems. Each section gives a detailed description of the blocks used and their operation within the overall print system. The final section describes the bi-lithic printhead construction and associated implications to the system due to its makeup.

4 Nomenclature

4.1 Bi-Lithic Printhead Notation

A bi-lithic based printhead is constructed from 2 printhead ICs of varying sizes. The notation M:N is used to express the size relationship of each IC, where M specifies one printhead IC in inches and N specifies the remaining printhead IC in inches.

The ‘SoPEC/MoPEC Bilithic Printhead Reference’ document [10] contains a description of the bilithic printhead and related terminology.

4.2 Definitions

The following terms are used throughout this specification:

-   Bi-lithic printhead Refers to printhead constructed from 2 printhead     ICs -   CPU Refers to CPU core, caching system and MMU. -   ISI-Bridge chip A device with a high speed interface (such as     USB2.0, Ethernet or IEEE1394) and one or more ISI interfaces. The     ISI-Bridge would be the ISIMaster for each of the ISI buses it     interfaces to. -   ISIMaster The ISIMaster is the only device allowed to initiate     communication on the Inter Sopec Interface (ISI) bus. The ISIMaster     interfaces with the host. -   ISISlave Multi-SoPEC systems will contain one or more ISISlave     SoPECs connected to the ISI bus. ISISlaves can only respond to     communication initiated by the ISIMaster. -   LEON Refers to the LEON CPU core. -   LineSyncMaster The LineSyncMaster device generates the line     synchronisation pulse that all SoPECs in the system must synchronise     their line outputs to. -   Multi-SoPEC Refers to SoPEC based print system with multiple SoPEC     devices -   Netpage Refers to page printed with tags (normally in infrared ink). -   PEC1 Refers to Print Engine Controller version 1, precursor to SoPEC     used to control printheads constructed from multiple angled     printhead segments. -   Printhead IC Single MEMS IC used to construct bi-lithic printhead -   PrintMaster The PrintMaster device is responsible for coordinating     all aspects of the print operation. There may only be one     PrintMaster in a system. -   QA Chip Quality Assurance Chip -   Storage SoPEC An ISISlave SoPEC used as a DRAM store and which does     not print. -   Tag Refers to pattern which encodes information about its position     and orientation which allow it to be optically located and its data     contents read.     4.3 Acronym and Abbreviations

The following acronyms and abbreviations are used in this specification

-   CFU Contone FIFO Unit -   CPU Central Processing Unit -   DIU DRAM Interface Unit -   DNC Dead Nozzle Compensator -   DRAM Dynamic Random Access Memory -   DWU DotLine Writer Unit -   GPIO General Purpose Input Output -   HCU Halftoner Compositor Unit -   ICU Interrupt Controller Unit -   ISI Inter SoPEC Interface -   LDB Lossless Bi-level Decoder -   LLU Line Loader Unit -   LSS Low Speed Serial interface -   MEMS Micro Electro Mechanical System -   MMU Memory Management Unit -   PCU SoPEC Controller Unit -   PHI PrintHead Interface -   PSS Power Save Storage Unit -   RDU Real-time Debug Unit -   ROM Read Only Memory -   SCB Serial Communication Block -   SFU Spot FIFO Unit -   SMG4 Silverbrook Modified Group 4. -   SoPEC Small office home office Print Engine Controller -   SRAM Static Random Access Memory -   TE Tag Encoder -   TFU Tag FIFO Unit -   TIM Timers Unit -   USB Universal Serial Bus     4.4 Pseudocode Notation

In general the pseudocode examples use C like statements with some exceptions. Symbol and naming convections used for pseudocode.

-   // Comment -   = Assignment -   ==,!=<,> Operator equal, not equal, less than, greater than -   +, −,*,/,% Operator addition, subtraction, multiply, divide, modulus -   &,|,^,<<,>>,˜ Bitwise AND, bitwise OR, bitwise exclusive OR, left     shift, right shift, complement -   AND,OR,NOT Logical AND, Logical OR, Logical inversion -   [XX:YY] Array/vector specifier -   {a, b, c} Concatenation operation -   ++, −− Increment and decrement     4.4.1 Register and Signal Naming Conventions

In general register naming uses the C style conventions with capitalization to denote word delimiters. Signals use RTL style notation where underscore denote word delimiters. There is a direct translation between both convention. For example the CmdSourceFifo register is equivalent to cmd_source_fifo signal.

4.5 State Machine Notation

State machines should be described using the pseudocode notation outlined above. State machine descriptions use the convention of underline to indicate the cause of a transition from one state to another and plain text (no underline) to indicate the effect of the transition i.e. signal transitions which occur when the new state is entered.

A sample state machine is shown in FIG. 1.

5 Printing Considerations

A bi-lithic printhead produces 1600 dpi bi-level dots. On low-diffusion paper, each ejected drop forms a 22.5 μm diameter dot. Dots are easily produced in isolation, allowing dispersed-dot dithering to be exploited to its fullest. Since the bi-lithic printhead is the width of the page and operates with a constant paper velocity, color planes are printed in perfect registration, allowing ideal dot-on-dot printing. Dot-on-dot printing minimizes ‘muddying’ of midtones caused by inter-color bleed. A page layout may contain a mixture of images, graphics and text. Continuous-tone (contone) images and graphics are reproduced using a stochastic dispersed-dot dither. Unlike a clustered-dot (or amplitude-modulated) dither, a dispersed-dot (or frequency-modulated) dither reproduces high spatial frequencies (i.e. image detail) almost to the limits of the dot resolution, while simultaneously reproducing lower spatial frequencies to their full color depth, when spatially integrated by the eye. A stochastic dither matrix is carefully designed to be free of objectionable low-frequency patterns when tiled across the image. As such its size typically exceeds the minimum size required to support a particular number of intensity levels (e.g. 16×16×8 bits for 257 intensity levels).

Human contrast sensitivity peaks at a spatial frequency of about 3 cycles per degree of visual field and then falls off logarithmically, decreasing by a factor of 100 beyond about 40 cycles per degree and becoming immeasurable beyond 60 cycles per degree [25][25]. At a normal viewing distance of 12 inches (about 300 mm), this translates roughly to 200-300 cycles per inch (cpi) on the printed page, or 400-600 samples per inch according to Nyquist's theorem.

In practice, contone resolution above about 300 ppi is of limited utility outside special applications such as medical imaging. Offset printing of magazines, for example, uses contone resolutions in the range 150 to 300 ppi. Higher resolutions contribute slightly to color error through the dither.

Black text and graphics are reproduced directly using bi-level black dots, and are therefore not anti-aliased (i.e. low-pass filtered) before being printed. Text should therefore be supersampled beyond the perceptual limits discussed above, to produce smoother edges when spatially integrated by the eye. Text resolution up to about 1200 dpi continues to contribute to perceived text sharpness (assuming low-diffusion paper, of course).

A Netpage printer, for example, may use a contone resolution of 267 ppi (i.e. 1600 dpi/6), and a black text and graphics resolution of 800 dpi. A high end office or departmental printer may use a contone resolution of 320 ppi (1600 dpi/5) and a black text and graphics resolution of 1600 dpi. Both formats are capable of exceeding the quality of commercial (offset) printing and photographic reproduction.

6 Document Data Flow

6.1 Considerations

Because of the page-width nature of the bi-lithic printhead, each page must be printed at a constant speed to avoid creating visible artifacts. This means that the printing speed can't be varied to match the input data rate. Document rasterization and document printing are therefore decoupled to ensure the printhead has a constant supply of data. A page is never printed until it is fully rasterized. This can be achieved by storing a compressed version of each rasterized page image in memory. This decoupling also allows the RIP(s) to run ahead of the printer when rasterizing simple pages, buying time to rasterize more complex pages.

Because contone color images are reproduced by stochastic dithering, but black text and line graphics are reproduced directly using dots, the compressed page image format contains a separate foreground bi-level black layer and background contone color layer. The black layer is composited over the contone layer after the contone layer is dithered (although the contone layer has an optional black component). A final layer of Netpage tags (in infrared or black ink) is optionally added to the page for printout.

FIG. 2 shows the flow of a document from computer system to printed page.

At 267 ppi for example, a A4 page (8.26 inches×11.7 inches) of contone CMYK data has a size of 26.3 MB. At 320 ppi, an A4 page of contone data has a size of 37.8 MB. Using lossy contone compression algorithms such as JPEG [27], contone images compress with a ratio up to 10:1 without noticeable loss of quality, giving compressed page sizes of 2.63 MB at 267 ppi and 3.78 MB at 320 ppi.

At 800 dpi, a A4 page of bi-level data has a size of 7.4 MB. At 1600 dpi, a Letter page of bi-level data has a size of 29.5 MB. Coherent data such as text compresses very well. Using lossless bi-level compression algorithms such as SMG4 fax as discussed in Section 8.1.2.3.1, ten-point plain text compresses with a ratio of about 50:1. Lossless bi-level compression across an average page is about 20:1 with 10:1 possible for pages which compress poorly. The requirement for SoPEC is to be able to print text at 10:1 compression. Assuming 10:1 compression gives compressed page sizes of 0.74 MB at 800 dpi, and 2.95 MB at 1600 dpi.

Once dithered, a page of CMYK contone image data consists of 116 MB of bi-level data. Using lossless bi-level compression algorithms on this data is pointless precisely because the optimal dither is stochastic—i.e. since it introduces hard-to-compress disorder.

Netpage tag data is optionally supplied with the page image. Rather than storing a compressed bi-level data layer for the Netpage tags, the tag data is stored in its raw form. Each tag is supplied up to 120 bits of raw variable data (combined with up to 56 bits of raw fixed data) and covers up to a 6 mm×6 mm area (at 1600 dpi). The absolute maximum number of tags on a A4 page is 15,540 when the tag is only 2 mm×2 mm (each tag is 126 dots×126 dots, for a total coverage of 148 tags×105 tags). 15,540 tags of 128 bits per tag gives a compressed tag page size of 0.24 MB.

The multi-layer compressed page image format therefore exploits the relative strengths of lossy JPEG contone image compression, lossless bi-level text compression, and tag encoding. The format is compact enough to be storage-efficient, and simple enough to allow straightforward real-time expansion during printing.

Since text and images normally don't overlap, the normal worst-case page image size is image only, while the normal best-case page image size is text only. The addition of worst case Netpage tags adds 0.24 MB to the page image size. The worst-case page image size is text over image plus tags. The average page size assumes a quarter of an average page contains images. Table 1 shows data sizes for compressed Letter page for these different options.

TABLE 1 Data sizes for A4 page (8.26 inches × 11.7 inches) 267 ppi contone 320 ppi contone 800 dpi bi-level 1600 dpi bi-level Image only (contone), 2.63 MB 3.78 MB 10:1 compression Text only (bi-level), 0.74 MB 2.95 MB 10:1 compression Netpage tags, 1600 dpi 0.24 MB 0.24 MB Worst case (text + 3.61 MB 6.67 MB image + tags) Average (text + 1.64 MB 4.25 MB 25% image + tags) 6.2 Document Data Flow

The Host PC rasterizes and compresses the incoming document on a page by page basis. The page is restructured into bands with one or more bands used to construct a page. The compressed data is then transferred to the SoPEC device via the USB link. A complete band is stored in SoPEC embedded memory. Once the band transfer is complete the SoPEC device reads the compressed data, expands the band, normalizes contone, bi-level and tag data to 1600 dpi and transfers the resultant calculated dots to the bi-lithic printhead.

The document data flow is

-   -   The RIP software rasterizes each page description and compress         the rasterized page image.     -   The infrared layer of the printed page optionally contains         encoded Netpage [5] tags at a programmable density.     -   The compressed page image is transferred to the SoPEC device via         the USB normally on a band by band basis.     -   The print engine takes the compressed page image and starts the         page expansion.     -   The first stage page expansion consists of 3 operations         performed in parallel     -   expansion of the JPEG-compressed contone layer     -   expansion of the SMG4 fax compressed bi-level layer     -   encoding and rendering of the bi-level tag data.     -   The second stage dithers the contone layer using a programmable         dither matrix, producing up to four bi-level layers at         full-resolution.     -   The second stage then composites the bi-level tag data layer,         the bi-level SMG4 fax de-compressed layer and up to four         bi-level JPEG de-compressed layers into the full-resolution page         image.     -   A fixative layer is also generated as required.     -   The last stage formats and prints the bi-level data through the         bi-lithic printhead via the printhead interface.

The SoPEC device can print a full resolution page with 6 color planes. Each of the color planes can be generated from compressed data through any channel (either JPEG compressed, bi-level SMG4 fax compressed, tag data generated, or fixative channel created) with a maximum number of 6 data channels from page RIP to bi-lithic printhead color planes.

The mapping of data channels to color planes is programmable, this allows for multiple color planes in the printhead to map to the same data channel to provide for redundancy in the printhead to assist dead nozzle compensation.

Also a data channel could be used to gate data from another data channel. For example in stencil mode, data from the bilevel data channel at 1600 dpi can be used to filter the contone data channel at 320 dpi, giving the effect of 1600 dpi contone image.

6.3 Page Considerations Due to SoPEC

The SoPEC device typically stores a complete page of document data on chip. The amount of storage available for compressed pages is limited to 2 Mbytes, imposing a fixed maximum on compressed page size. A comparison of the compressed image sizes in Table 2 indicates that SoPEC would not be capable of printing worst case pages unless they are split into bands and printing commences before all the bands for the page have been downloaded. The page sizes in the table are shown for comparison purposes and would be considered reasonable for a professional level printing system. The SoPEC device is aimed at the consumer level and would not be required to print pages of that complexity. Target document types for the SoPEC device are shown Table 2.

TABLE 2 Page content targets for SoPEC Size Page Content Description Calculation (MByte) Best Case picture Image, 8.26 × 11.7 × 267 × 1.97 267 ppi with 3 colors, 267 × 3 @ 10:1 A4 size Full page text, 800 dpi 8.26 × 11.7 × 800 × 0.74 A4 size 800 @ 10:1 Mixed Graphics and Text 6 × 4 × 267 × 267 × 1.55 Image of 6 inches × 4 3 @ 5:1 inches @ 267 ppi and 800 × 800 × 73 @ 10:1 3 colors Remaining area text ~73 inches², 800 dpi Best Case Photo, 3 Colors, 6.6 Mpixel @ 10:1 2.00 6.6 MegaPixel Image

If a document with more complex pages is required, the page RIP software in the host PC can determine that there is insufficient memory storage in the SoPEC for that document. In such cases the RIP software can take two courses of action. It can increase the compression ratio until the compressed page size will fit in the SoPEC device, at the expense of document quality, or divide the page into bands and allow SoPEC to begin printing a page band before all bands for that page are downloaded. Once SoPEC starts printing a page it cannot stop, if SoPEC consumes compressed data faster than the bands can be downloaded a buffer underrun error could occur causing the print to fail. A buffer underrun occurs if a line synchronisation pulse is received before a line of data has been transferred to the printhead.

Other options which can be considered if the page does not fit completely into the compressed page store are to slow the printing or to use multiple SoPECs to print parts of the page. A Storage SoPEC (Section 7.2.5) could be added to the system to provide guaranteed bandwidth data delivery. The print system could also be constructed using an ISI-Bridge chip (Section 7.2.6) to provide guaranteed data delivery.

7 Memjet Printer Architecture

The SoPEC device can be used in several printer configurations and architectures.

In the general sense every SoPEC based printer architecture will contain:

-   -   One or more SoPEC devices.     -   One or more bi-lithic printheads.     -   Two or more LSS busses.     -   Two or more QA chips.     -   USB 1.1 connection to host or ISI connection to Bridge Chip.     -   ISI bus connection between SoPECs (when multiple SoPECs are         used).

Some example printer configurations as outlined in Section 7.2. The various system components are outlined briefly in Section 7.1.

7.1 System Components

7.1.1 SoPEC Print Engine Controller

The SoPEC device contains several system on a chip (SoC) components, as well as the print engine pipeline control application specific logic.

7.1.1.1 Print Engine Pipeline (PEP) Logic

The PEP reads compressed page store data from the embedded memory, optionally decompresses the data and formats it for sending to the printhead. The print engine pipeline functionality includes expanding the page image, dithering the contone layer, compositing the black layer over the contone layer, rendering of Netpage tags, compensation for dead nozzles in the printhead, and sending the resultant image to the bi-lithic printhead.

7.1.1.2 Embedded CPU

SoPEC contains an embedded CPU for general purpose system configuration and management. The CPU performs page and band header processing, motor control and sensor monitoring (via the GPIO) and other system control functions. The CPU can perform buffer management or report buffer status to the host. The CPU can optionally run vendor application specific code for general print control such as paper ready monitoring and LED status update.

7.1.1.3 Embedded Memory Buffer

A 2.5 Mbyte embedded memory buffer is integrated onto the SoPEC device, of which approximately 2 Mbytes are available for compressed page store data. A compressed page is divided into one or more bands, with a number of bands stored in memory. As a band of the page is consumed by the PEP for printing a new band can be downloaded. The new band may be for the current page or the next page.

Using banding it is possible to begin printing a page before the complete compressed page is downloaded, but care must be taken to ensure that data is always available for printing or a buffer underrun may occur.

An Storage SoPEC acting as a memory buffer (Section 7.2.5) or an ISI-Bridge chip with attached DRAM (Section 7.2.6) could be used to provide guaranteed data delivery.

7.1.1.4 Embedded USB 1.1 Device

The embedded USB 1.1 device accepts compressed page data and control commands from the host PC, and facilitates the data transfer to either embedded memory or to another SoPEC device in multi-SoPEC systems.

7.1.2 Bi-lithic Printhead

The printhead is constructed by abutting 2 printhead ICs together. The printhead ICs can vary in size from 2 inches to 8 inches, so to produce an A4 printhead several combinations are possible. For example two printhead ICs of 7 inches and 3 inches could be used to create a A4 printhead (the notation is 7:3). Similarly 6 and 4 combination (6:4), or 5:5 combination. For an A3 printhead it can be constructed from 8:6 or an 7:7 printhead IC combination. For photographic printing smaller printheads can be constructed.

7.1.3 LSS Interface Bus

Each SoPEC device has 2 LSS system buses for communication with QA devices for system authentication and ink usage accounting. The number of QA devices per bus and their position in the system is unrestricted with the exception that PRINTER_QA and INK_QA devices should be on separate LSS busses.

7.1.4 QA Devices

Each SoPEC system can have several QA devices. Normally each printing SoPEC will have an associated PRINTER_QA. Ink cartridges will contain an INK_QA chip. PRINTER_QA and INK_QA devices should be on separate LSS busses. All QA chips in the system are physically identical with flash memory contents defining PRINTER_QA from INK_QA chip.

7.1.5 ISI Interface

The Inter-SoPEC Interface (ISI) provides a communication channel between SoPECs in a multi-SoPEC system. The ISIMaster can be SoPEC device or an ISI-Bridge chip depending on the printer configuration. Both compressed data and control commands are transferred via the interface.

7.1.6 ISI-Bridge Chip

A device, other than a SoPEC with a USB connection, which provides print data to a number of slave SoPECs. A bridge chip will typically have a high bandwidth connection, such as USB2.0, Ethernet or IEEE1394, to a host and may have an attached external DRAM for compressed page storage. A bridge chip would have one or more ISI interfaces. The use of multiple ISI buses would allow the construction of independent print systems within the one printer. The ISI-Bridge would be the ISIMaster for each of the ISI buses it interfaces to.

7.2 Possible SoPEC Systems

Several possible SoPEC based system architectures exist. The following sections outline some possible architectures. It is possible to have extra SoPEC devices in the system used for DRAM storage. The QA chip configurations shown are indicative of the flexibility of LSS bus architecture, but not limited to those configurations.

7.2.1 A4 Simplex with 1 SoPEC Device

In FIG. 3, a single SoPEC device can be used to control two printhead ICs. The SoPEC receives compressed data through the USB device from the host. The compressed data is processed and transferred to the printhead.

7.2.2 A4 Duplex with 2 SoPEC Devices

In FIG. 4, two SoPEC devices are used to control two bi-lithic printheads, each with two printhead ICs. Each bi-lithic printhead prints to opposite sides of the same page to achieve duplex printing. The SoPEC connected to the USB is the ISIMaster SoPEC, the remaining SoPEC is an ISISlave. The ISIMaster receives all the compressed page data for both SoPECs and re-distributes the compressed data over the Inter-SoPEC Interface (ISI) bus.

It may not be possible to print an A4 page every 2 seconds in this configuration since the USB 1.1 connection to the host may not have enough bandwidth. An alternative would be for each SoPEC to have its own USB 1.1 connection. This would allow a faster average print speed.

7.2.3 A3 Simplex with 2 SoPEC Devices

In FIG. 5, two SoPEC devices are used to control one A3 bi-lithic printhead. Each SoPEC controls only one printhead IC (the remaining PHI port typically remains idle). This system uses the SoPEC with the USB connection as the ISIMaster. In this dual SoPEC configuration the compressed page store data is split across 2 SoPECs giving a total of 4 Mbyte page store, this allows the system to use compression rates as in an A4 architecture, but with the increased page size of A3. The ISIMaster receives all the compressed page data for all SoPECs and re-distributes the compressed data over the Inter-SoPEC Interface (ISI) bus.

It may not be possible to print an A3 page every 2 seconds in this configuration since the USB 1.1 connection to the host will only have enough bandwidth to supply 2 Mbytes every 2 seconds. Pages which require more than 2 MBytes every 2 seconds will therefore print more slowly. An alternative would be for each SoPEC to have its own USB 1.1 connection. This would allow a faster average print speed.

7.2.4 A3 Duplex with 4 SoPEC Devices

In FIG. 6 a 4 SoPEC system is shown. It contains 2 A3 bi-lithic printheads, one for each side of an A3 page. Each printhead contain 2 printhead ICs, each printhead IC is controlled by an independent SoPEC device, with the remaining PHI port typically unused. Again the SoPEC with USB 1.1 connection is the ISIMaster with the other SoPECs as ISISlaves. In total, the system contains 8 Mbytes of compressed page store (2 Mbytes per SoPEC), so the increased page size does not degrade the system print quality, from that of an A4 simplex printer. The ISIMaster receives all the compressed page data for all SoPECs and re-distributes the compressed data over the Inter-SoPEC Interface (ISI) bus.

It may not be possible to print an A3 page every 2 seconds in this configuration since the USB 1.1 connection to the host will only have enough bandwidth to supply 2 Mbytes every 2 seconds. Pages which require more than 2 MBytes every 2 seconds will therefore print more slowly. An alternative would be for each SoPEC or set of SoPECs on the same side of the page to have their own USB 1.1 connection (as ISISlaves may also have direct USB connections to the host). This would allow a faster average print speed.

7.2.5 SoPEC DRAM storage solution: A4 Simplex with 1 printing SoPEC and 1 memory SoPEC Extra SoPECs can be used for DRAM storage e.g. in FIG. 7 an A4 simplex printer can be built with a single extra SoPEC used for DRAM storage. The DRAM SoPEC can provide guaranteed bandwidth delivery of data to the printing SoPEC. SoPEC configurations can have multiple extra SoPECs used for DRAM storage.

7.2.6 ISI-Bridge Chip Solution: A3 Duplex System with 4 SoPEC Devices

In FIG. 8, an ISI-Bridge chip provides slave-only ISI connections to SoPEC devices. FIG. 8 shows a ISI-Bridge chip with 2 separate ISI ports. The ISI-Bridge chip is the ISIMaster on each of the ISI busses it is connected to. All connected SoPECs are ISISlaves. The ISI-Bridge chip will typically have a high bandwidth connection to a host and may have an attached external DRAM for compressed page storage.

An alternative to having a ISI-Bridge chip would be for each SoPEC or each set of SoPECs on the same side of a page to have their own USB 1.1 connection. This would allow a faster average print speed.

8 Page Format and Printflow

When rendering a page, the RIP produces a page header and a number of bands (a non-blank page requires at least one band) for a page. The page header contains high level rendering parameters, and each band contains compressed page data. The size of the band will depend on the memory available to the RIP, the speed of the RIP, and the amount of memory remaining in SoPEC while printing the previous band(s). FIG. 9 shows the high level data structure of a number of pages with different numbers of bands in the page.

Each compressed band contains a mandatory band header, an optional bi-level plane, optional sets of interleaved contone planes, and an optional tag data plane (for Netpage enabled applications). Since each of these planes is optional¹, the band header specifies which planes are included with the band. FIG. 10 gives a high-level breakdown of the contents of a page band. ¹Although a band must contain at least one plane

A single SoPEC has maximum rendering restrictions as follows:

-   -   1 bi-level plane     -   1 contone interleaved plane set containing a maximum of 4         contone planes     -   1 tag data plane     -   a bi-lithic printhead with a maximum of 2 printhead ICs

The requirement for single-sided A4 single SoPEC printing is

-   -   average contone JPEG compression ratio of 10:1, with a local         minimum compression ratio of 5:1 for a single line of         interleaved JPEG blocks.     -   average bi-level compression ratio of 10:1, with a local minimum         compression ratio of 1:1 for a single line.

If the page contains rendering parameters that exceed these specifications, then the RIP or the Host PC must split the page into a format that can be handled by a single SoPEC.

In the general case, the SoPEC CPU must analyze the page and band headers and generate an appropriate set of register write commands to configure the units in SoPEC for that page. The various bands are passed to the destination SoPEC(s) to locations in DRAM determined by the host.

The host keeps a memory map for the DRAM, and ensures that as a band is passed to a SoPEC, it is stored in a suitable free area in DRAM. Each SoPEC is connected to the ISI bus or USB bus via its Serial communication Block (SCB). The SoPEC CPU configures the SCB to allow compressed data bands to pass from the USB or ISI through the SCB to SoPEC DRAM. FIG. 11 shows an example data flow for a page destined to be printed by a single SoPEC. Band usage information is generated by the individual SoPECs and passed back to the host.

SoPEC has an addressing mechanism that permits circular band memory allocation, thus facilitating easy memory management. However it is not strictly necessary that all bands be stored together. As long as the appropriate registers in SoPEC are set up for each band, and a given band is contiguous², the memory can be allocated in any way. ²Contiguous allocation also includes wrapping around in SoPEC's band store memory.

8.1 Print Engine Example Page Format

This section describes a possible format of compressed pages expected by the embedded CPU in SoPEC. The format is generated by software in the host PC and interpreted by embedded software in SoPEC. This section indicates the type of information in a page format structure, but implementations need not be limited to this format. The host PC can optionally perform the majority of the header processing.

The compressed format and the print engines are designed to allow real-time page expansion during printing, to ensure that printing is never interrupted in the middle of a page due to data underrun.

The page format described here is for a single black bi-level layer, a contone layer, and a Netpage tag layer. The black bi-level layer is defined to composite over the contone layer.

The black bi-level layer consists of a bitmap containing a 1-bit opacity for each pixel. This black layer matte has a resolution which is an integer or non-integer factor of the printer's dot resolution.

The highest supported resolution is 1600 dpi, i.e. the printer's full dot resolution.

The contone layer, optionally passed in as YCrCb, consists of a 24-bit CMY or 32-bit CMYK color for each pixel. This contone image has a resolution which is an integer or non-integer factor of the printer's dot resolution. The requirement for a single SoPEC is to support 1 side per 2 seconds A4/Letter printing at a resolution of 267 ppi, i.e. one-sixth the printer's dot resolution.

Non-integer scaling can be performed on both the contone and bi-level images. Only integer scaling can be performed on the tag data.

The black bi-level layer and the contone layer are both in compressed form for efficient storage in the printer's internal memory.

8.1.1 Page Structure

A single SoPEC is able to print with full edge bleed for Letter and A3 via different stitch part combinations of the bi-lithic printhead. It imposes no margins and so has a printable page area which corresponds to the size of its paper. The target page size is constrained by the printable page area, less the explicit (target) left and top margins specified in the page description. These relationships are illustrated below.

8.1.2 Compressed Page Format

Apart from being implicitly defined in relation to the printable page area, each page description is complete and self-contained. There is no data stored separately from the page description to which the page description refers.³ The page description consists of a page header which describes the size and resolution of the page, followed by one or more page bands which describe the actual page content. ³SoPEC relies on dither matrices and tag structures to have already been set up, but these are not considered to be part of a general page format. It is trivial to extend the page format to allow exact specification of dither matrices and tag structures.

8.1.2.1 Page Header

Table 3 shows an example format of a page header.

TABLE 3 Page header format field format description signature 16-bit Page header format signature. integer version 16-bit Page header format version number. integer structure size 16-bit Size of page header. integer band count 16-bit Number of bands specified for integer this page. target resolution 16-bit Resolution of target page. This (dpi) integer is always 1600 for the Memjet printer. target page width 16-bit Width of target page, in dots. integer target page height 32-bit Height of target page, in dots. integer target left margin 16-bit Width of target left margin, in for black and integer dots, for black and contone. contone target top margin 16-bit Height of target top margin, in for black and integer dots, for black and contone. contone target right 16-bit Width of target right margin, in margin for black integer dots, for black and contone. and contone target bottom 16-bit Height of target bottom margin, margin for black integer in dots, for black and contone. and contone target left 16-bit Width of target left margin, in margin for tags integer dots, for tags. target top 16-bit Height of target top margin, in margin for tags integer dots, for tags. target right 16-bit Width of target right margin, margin for tags integer in dots, for tags. target bottom 16-bit Height of target bottom margin, margin for tags integer in dots, for tags. generate tags 16-bit Specifies whether to generate integer tags for this page (0 - no, 1 - yes). fixed tag data 128-bit This is only valid if generate integer tags is set. tag vertical 16-bit Scale factor in vertical direc- scale factor integer tion from tag data resolution to target resolution. Valid range = 1-511. Integer scaling only tag horizontal 16-bit Scale factor in horizontal scale factor integer direction from tag data reso- lution to target resolution. Valid range = 1-511. Integer scaling only. bi-level layer 16-bit Scale factor in vertical vertical scale factor integer direction from bi-level resolution to target reso- lution (must be 1 or greater). May be non-integer. Expressed as a fraction with upper 8-bits the numerator and the lower 8 bits the denominator. bi-level layer 16-bit Scale factor in horizontal horizontal integer direction from bi-level reso- scale factor lution to target resolution (must be 1 or greater). May be non-integer. Expressed as a fraction with upper 8-bits the numerator and the lower 8 bits the denominator. bi-level layer 16-bit Width of bi-level layer page, page width integer in pixels. bi-level layer 32-bit Height of bi-level layer page, page height integer in pixels. contone flags 16 bit Defines the color conversion integer that is required for the JPEG data. Bits 2-0 specify how many contone planes there are (e.g. 3 for CMY and 4 for CMYK). Bit 3 specifies whether the first 3 color planes need to be converted back from YCrCb to CMY. Only valid if b2-0 = 3 or 4. 0 - no conversion, leave JPEG colors alone 1 - color convert. Bits 7-4 specifies whether the YCrCb was generated directly from CMY, or whether it was converted to RGB first via the step: R = 255-C, G = 255-M, B = 255-Y. Each of the color planes can be individually inverted. Bit 4: 0 - do not invert color plane 0 1 - invert color plane 0 Bit 5: 0 - do not invert color plane 1 1 - invert color plane 1 Bit 6: 0 - do not invert color plane 2 1 - invert color plane 2 Bit 7: 0 - do not invert color plane 3 1 - invert color plane 3 Bit 8 specifies whether the con- tone data is JPEG compressed or non-compressed: 0 - JPEG compressed 1 - non-compressed The remaining bits are reserved (0). contone vertical 16-bit Scale factor in vertical direction scale factor integer from contone channel resolution to target resolution. Valid range = 1-255. May be non-integer. Expressed as a fraction with upper 8-bits the numerator and the lower 8 bits the denominator. contone 16-bit Scale factor in horizontal direction horizontal integer from contone channel resolution to scale factor target resolution. Valid range = 1-255. May be non-integer. Expressed as a fraction with upper 8-bits the numerator and the lower 8 bits the denominator. contone page 16-bit Width of contone page, in contone width integer pixels. contone page 32-bit Height of contone page, in contone height integer pixels. reserved up to 128 Reserved and 0 pads out page header bytes to multiple of 128 bytes.

The page header contains a signature and version which allow the CPU to identify the page header format. If the signature and/or version are missing or incompatible with the CPU, then the CPU can reject the page.

The contone flags define how many contone layers are present, which typically is used for defining whether the contone layer is CMY or CMYK. Additionally, if the color planes are CMY, they can be optionally stored as YCrCb, and further optionally color space converted from CMY directly or via RGB. Finally the contone data is specified as being either JPEG compressed or non-compressed. The page header defines the resolution and size of the target page. The bi-level and contone layers are clipped to the target page if necessary. This happens whenever the bi-level or contone scale factors are not factors of the target page width or height.

The target left, top, right and bottom margins define the positioning of the target page within the printable page area.

The tag parameters specify whether or not Netpage tags should be produced for this page and what orientation the tags should be produced at (landscape or portrait mode). The fixed tag data is also provided.

The contone, bi-level and tag layer parameters define the page size and the scale factors.

8.1.2.2 Band Format

Table 4 shows the format of the page band header.

TABLE 4 Band header format field format description signature 16-bit Page band header format signature. integer version 16-bit Page band header format version integer number. structure size 16-bit Size of page band header. integer bi-level layer 16-bit Height of bi-level layer band, in band height integer black pixels. bi-level layer 32-bit Size of bi-level layer band data, band data size integer in bytes. contone band 16-bit Height of contone band, in contone height integer pixels. contone band 32-bit Size of contone plane band data, data size integer in bytes. tag band 16-bit Height of tag band, in dots. height integer tag band 32-bit Size of unencoded tag data band, data size integer in bytes. Can be 0 which indicates that no tag data is provided. reserved up to 128 Reserved and 0 pads out band bytes header to multiple of 128 bytes.

The bi-level layer parameters define the height of the black band, and the size of its compressed band data. The variable-size black data follows the page band header.

The contone layer parameters define the height of the contone band, and the size of its compressed page data. The variable-size contone data follows the black data.

The tag band data is the set of variable tag data half-lines as required by the tag encoder. The format of the tag data is found in Section 26.5.2. The tag band data follows the contone data.

Table 5 shows the format of the variable-size compressed band data which follows the page band header.

TABLE 5 Page band data format field format Description black data Modified G4 Compressed bi-level facsimile bitstream⁴ layer. contone data JPEG bytestream Compressed contone datalayer. tag data map Tag data array Tag data format. See Section 26.5.2. ⁴ See section 8.1.2.3 on page 36 for note regarding the use of this standard

The start of each variable-size segment of band data should be aligned to a 256-bit DRAM word boundary.

The following sections describe the format of the compressed bi-level layers and the compressed contone layer. section 26.5.1 on page 410 describes the format of the tag data structures.

8.1.2.3 Bi-Level Data Compression

The (typically 1600 dpi) black bi-level layer is losslessly compressed using Silverbrook Modified Group 4 (SMG4) compression which is a version of Group 4 Facsimile compression [22] without Huffman and with simplified run length encodings. Typically compression ratIOs exceed 10:1. The encoding are listed in Table 6 and Table 7. ⁴ See section 8.123 on page 36 for note regarding the use of this standard

TABLE 6 Bi-Level group 4 facsimile style compression encodings Encoding Description same as Group 4 1000 Pass Command: a0

Facsimile b2, skip next two edges 1 Vertical(0): a0

b1, color = !color 110 Vertical(1): a0

b1 + 1, color = !color 010 Vertical(−1): a0

b1 − 1, color = !color 110000 Vertical(2): a0

b1 + 2, color = !color 010000 Vertical(−2): a0

b1 − 2, color = !color Unique to this 100000 Vertical(3): a0

implementation b1 + 3, color = !color 000000 Vertical(−3): a0

b1 − 3, color = !color <RL><RL>100 Horizontal: a0

a0 + <RL> + <RL>

SMG4 has a pass through mode to cope with local negative compression. Pass through mode is activated by a special run-length code. Pass through mode continues to either end of line or for a pre-programmed number of bits, whichever is shorter. The special run-length code is always executed as a run-length code, followed by pass through. The pass through escape code is a medium length run-length with a run of less than or equal to 31.

TABLE 7 Run length (RL) encodings Encoding Description Unique to this RRRRR1 Short Black Runlength implementation (5 bits) RRRRR1 Short White Runlength (5 bits) RRRRRRRRRR10 Medium Black Runlength (10 bits) RRRRRRRR10 Medium White Runlength (8 bits) RRRRRRRRRR10 Medium Black Runlength with RRRRRRRRRR <= 31, Enter pass through RRRRRRRR10 Medium White Runlength with RRRRRRRR <= 31, Enter pass through RRRRRRRRRRRRRRR00 Long Black Runlength (15 bits) RRRRRRRRRRRRRRR00 Long White Runlength (15 bits)

Since the compression is a bitstream, the encodings are read right (least significant bit) to left (most significant bit). The run lengths given as RRRR in Table are read in the same way (least significant bit at the right to most significant bit at the left).

Each band of bi-level data is optionally self contained. The first line of each band therefore is based on a ‘previous’ blank line or the last line of the previous band.

8.1.2.3.1 Group 3 and 4 Facsimile Compression

The Group 3 Facsimile compression algorithm [22] losslessly compresses bi-level data for transmission over slow and noisy and noisy telephone lines. The bi-level data represents scanned black text and graphics on a while background, and the algorithm is tuned for this class of images (it is explicitly not tuned, for example, for halftoned bi-level images). The 1D Group 3 algorithm runlength-encodes each scanline and then Huffman-encodes the resulting runlengths. Runlengths in the range 0 to 63 are coded with terminating codes. Runlengths in the range 64 to 2623 are coded with make-up codes, each representing a multiple of 64, followed by a terminating code. Runlengths exceeding 2623 are coded with multiple make-up codes followed by a terminating code. The Huffman tables are fixed, but are separately tuned for black and white runs (except for make-up codes above 1728, which are common). When possible, the 2D Group 3 algorithm encodes a scanline as a set of short edge deltas (0, +1, +2, +3) with reference to the previous scanline. The delta symbols are entropy-encoded (so that the zero delta symbol is only one bit long etc.) Edges within a 2D-encoded line which can't be delta-encoded are runlength-encoded, and are identified by a prefix. 1D- and 2D-encoded lines are marked differently. 1D-encoded lines are generated at regular intervals, whether actually required or not, to ensure that the decoder can recover from line noise with minimal image degradation. 2D Group 3 achieves compression ratIOs of up to 6:1 [32]. The Group 4 Facsimile algorithm [22] losslessly compresses bi-level data for transmission over error-free communications lines (i.e. the lines are truly error-free, or error-correction is done at a lower protocol level). The Group 4 algorithm is based on the 2D Group 3 algorithm, with the essential modification that since transmission is assumed to be error-free, 1D-encoded lines are no longer generated at regular intervals as an aid to error-recovery. Group 4 achieves compression ratIOs ranging from 20:1 to 60:1 for the CCITT set of test images [32].

The design goals and performance of the Group 4 compression algorithm qualify it as a compression algorithm for the bi-level layers. However, its Huffman tables are tuned to a lower scanning resolution (100-400 dpi), and it encodes runlengths exceeding 2623 awkwardly.

8.1.2.4 Contone Data Compression

The contone layer (CMYK) is either a non-compressed bytestream or is compressed to an interleaved JPEG bytestream. The JPEG bytestream is complete and self-contained. It contains all data required for decompression, including quantization and Huffman tables.

The contone data is optionally converted to YCrCb before being compressed (there is no specific advantage in color-space converting if not compressing). Additionally, the CMY contone pixels are optionally converted (on an individual basis) to RGB before color conversion using R=255-C, G=255-M, B=255-Y. Optional bitwise inversion of the K plane may also be performed. Note that this CMY to RGB conversion is not intended to be accurate for display purposes, but rather for the purposes of later converting to YCrCb. The inverse transform will be applied before printing.

8.1.2.4.1 JPEG Compression

The JPEG compression algorithm [27] lossily compresses a contone image at a specified quality level. It introduces imperceptible image degradation at compression ratIOs below 5:1, and negligible image degradation at compression ratIOs below 10:1 [33].

JPEG typically first transforms the image into a color space which separates luminance and chrominance into separate color channels. This allows the chrominance channels to be subsampled without appreciable loss because of the human visual system's relatively greater sensitivity to luminance than chrominance. After this first step, each color channel is compressed separately.

The image is divided into 8×8 pixel blocks. Each block is then transformed into the frequency domain via a discrete cosine transform (DCT). This transformation has the effect of concentrating image energy in relatively lower-frequency coefficients, which allows higher-frequency coefficients to be more crudely quantized. This quantization is the principal source of compression in JPEG. Further compression is achieved by ordering coefficients by frequency to maximize the likelihood of adjacent zero coefficients, and then runlength-encoding runs of zeroes. Finally, the runlengths and non-zero frequency coefficients are entropy coded. Decompression is the inverse process of compression.

8.1.2.4.2 Non-Compressed Format

If the contone data is non-compressed, it must be in a block-based format bytestream with the same pixel order as would be produced by a JPEG decoder. The bytestream therefore consists of a series of 8×8 block of the original image, starting with the top left 8×8 block, and working horizontally across the page (as it will be printed) until the top rightmost 8×8 block, then the next row of 8×8 blocks (left to right) and so on until the lower row of 8×8 blocks (left to right). Each 8×8 block consists of 64 8-bit pixels for color plane 0 (representing 8 rows of 8 pixels in the order top left to bottom right) followed by 64 8-bit pixels for color plane 1 and so on for up to a maximum of 4 color planes.

If the original image is not a multiple of 8 pixels in X or Y, padding must be present (the extra pixel data will be ignored by the setting of margins).

8.1.2.4.3 Compressed Format

If the contone data is compressed the first memory band contains JPEG headers (including tables) plus MCUs (minimum coded units). The ratio of space between the various color planes in the JPEG stream is 1:1:1:1. No subsampling is permitted. Banding can be completely arbitrary i.e there can be multiple JPEG images per band or 1 JPEG image divided over multiple bands. The break between bands is only memory alignment based.

8.1.2.4.4 Conversion of RGB to YCrCb (in RIP)

YCrCb is defined as per CCIR 601-1 [24] except that Y, Cr and Cb are normalized to occupy all 256 levels of an 8-bit binary encoding and take account of the actual hardware implementation of the inverse transform within SoPEC.

The exact color conversion computation is as follows:

-   -   Y*=(9805/32768)R+(19235/32768)G+(3728/32768)B     -   Cr*=(16375/32768)R−(13716/32768)G−(2659/32768)B+128     -   Cb*=−(5529/32768)R−(10846/32768)G+(16375/32768)B+128

Y, Cr and Cb are obtained by rounding to the nearest integer. There is no need for saturation since ranges of Y*, Cr* and Cb* after rounding are [0-255], [1-255] and [1-255] respectively. Note that full accuracy is possible with 24 bits. See [14] for more information.

SoPEC ASIC

9 Overview

The Small Office Home Office Print Engine Controller (SoPEC) is a page rendering engine ASIC that takes compressed page images as input, and produces decompressed page images at up to 6 channels of bi-level dot data as output. The bi-level dot data is generated for the Memjet bi-lithic printhead. The dot generation process takes account of printhead construction, dead nozzles, and allows for fixative generation.

A single SoPEC can control 2 bi-lithic printheads and up to 6 color channels at 10,000 lines/sec⁵, equating to 30 pages per minute. A single SoPEC can perform full-bleed printing of A3, A4 and Letter pages. The 6 channels of colored ink are the expected maximum in a consumer SOHO, or office Bi-lithic printing environment: ⁵10,000 lines per second equates to 30 A4/Letter pages per minute at 1600 dpi

-   -   CMY, for regular color printing.     -   K, for black text, line graphics and gray-scale printing.     -   IR (infrared), for Netpage-enabled [5] applications.     -   F (fixative), to enable printing at high speed. Because the         bi-lithic printer is capable of printing so fast, a fixative may         be required to enable the ink to dry before the page touches the         page already printed. Otherwise the pages may bleed on each         other. In low speed printing environments the fixative may not         be required.

SoPEC is color space agnostic. Although it can accept contone data as CMYX or RGBX, where X is an optional 4th channel, it also can accept contone data in any print color space. Additionally, SoPEC provides a mechanism for arbitrary mapping of input channels to output channels, including combining dots for ink optimization, generation of channels based on any number of other channels etc. However, inputs are typically CMYK for contone input, K for the bi-level input, and the optional Netpage tag dots are typically rendered to an infra-red layer. A fixative channel is typically generated for fast printing applications.

SoPEC is resolution agnostic. It merely provides a mapping between input resolutions and output resolutions by means of scale factors. The expected output resolution is 1600 dpi, but SoPEC actually has no knowledge of the physical resolution of the Bi-lithic printhead.

SoPEC is page-length agnostic. Successive pages are typically split into bands and downloaded into the page store as each band of information is consumed and becomes free.

SoPEC provides an interface for synchronization with other SoPECs. This allows simple multi-SoPEC solutions for simultaneous A3/A4/Letter duplex printing. However, SoPEC is also capable of printing only a portion of a page image. Combining synchronization functionality with partial page rendering allows multiple SoPECs to be readily combined for alternative printing requirements including simultaneous duplex printing and wide format printing.

Table 8 lists some of the features and corresponding benefits of SoPEC.

TABLE 8 Features and Benefits of SoPEC Feature Benefits Optimised print 30 ppm full page photographic architecture in quality color printing from a hardware desktop PC 0.13 micron CMOS High speed (>3 million Low cost transistors) High functionality 900 Million dots Extremely fast page generation per second 10,000 lines per 0.5 A4/Letter pages per SoPEC second at 1600 dpi chip per second 1 chip drives up to Low cost page-width printers 133,920 nozzles 1 chip drives up to 6 99% of SoHo printers can use color planes 1 SoPEC device Integrated DRAM No external memory required, leading to low cost systems Power saving SoPEC can enter a power saving sleep mode sleep mode to reduce power dissipation between print jobs JPEG expansion Low bandwidth from PC Low memory requirements in printer Lossless bitplane High resolution text and line expansion art with low bandwidth from PC (e.g. over USB) Netpage tag expansion Generates interactive paper Stochastic dispersed Optically smooth image quality dot dither No moire effects Hardware compositor Pages composited in real-time for 6 image planes Dead nozzle compensation Extends printhead life and yield Reduces printhead cost Color space agnostic Compatible with all inksets and image sources including RGB, CMYK, spot, CIE L*a*b*, hexachrome, YCrCbK, sRGB and other Color space conversion Higher quality/lower bandwidth Computer interface USB1.1 interface to host and ISI interface to ISI-Bridge chip thereby allowing connection to IEEE 1394, Bluetooth etc. Cascadable in resolution Printers of any resolution Cascadable in color depth Special color sets e.g. hexachrome can be used Cascadable in image size Printers of any width up to 16 inches Cascadable in pages Printers can print both sides simultaneously Cascadable in speed Higher speeds are possible by having each SoPEC print one vertical strip of the page. Fixative channel Extremely fast ink drying data generation without wastage Built-in security Revenue models are protected Undercolor removal on Reduced ink usage dot-by-dot basis Does not require fonts for No font substitution or high speed operation missing fonts Flexible printhead Many configurations of configuration printheads are supported by one chip type Drives Bi-lithic No print driver chips required, printheads directly results in lower cost Determines dot accurate Removes need for physical ink ink usage monitoring system in ink cartridges 9.1 Printing Rates

The required printing rate for SoPEC is 30 sheets per minute with an inter-sheet spacing of 4 cm. To achieve a 30 sheets per minute print rate, this requires:

-   -   300 mm×63 (dot/mm)/2 sec=105.8 μseconds per line, with no         inter-sheet gap.     -   340 mm×63 (dot/mm)/2 sec=93.3 μseconds per line, with a 4 cm         inter-sheet gap.

A printline for an A4 page consists of 13824 nozzles across the page [2]. At a system clock rate of 160 MHz 13824 dots of data can be generated in 86.4 μseconds. Therefore data can be generated fast enough to meet the printing speed requirement. It is necessary to deliver this print data to the print-heads.

Printheads can be made up of 5:5, 6:4, 7:3 and 8:2 inch printhead combinations [2]. Print data is transferred to both print heads in a pair simultaneously. This means the longest time to print a line is determined by the time to transfer print data to the longest print segment. There are 9744 nozzles across a 7 inch printhead. The print data is transferred to the printhead at a rate of 106 MHz (2/3 of the system clock rate) per color plane. This means that it will take 91.9 μs to transfer a single line for a 7:3 printhead configuration. So we can meet the requirement of 30 sheets per minute printing with a 4 cm gap with a 7:3 printhead combination. There are 11160 across an 8 inch printhead. To transfer the data to the printhead at 106 MHz will take 105.3 μs. So an 8:2 printhead combination printing with an inter-sheet gap will print slower than 30 sheets per minute.

9.2 SoPEC Basic Architecture

From the highest point of view the SoPEC device consists of 3 distinct subsystems

-   -   CPU Subsystem     -   DRAM Subsystem     -   Print Engine Pipeline (PEP) Subsystem

See FIG. 13 for a block level diagram of SoPEC.

9.2.1 CPU Subsystem

The CPU subsystem controls and configures all aspects of the other subsystems. It provides general support for interfacing and synchronising the external printer with the internal print engine. It also controls the low speed communication to the QA chips. The CPU subsystem contains various peripherals to aid the CPU, such as GPIO (includes motor control), interrupt controller, LSS Master and general timers. The Serial Communications Block (SCB) on the CPU subsystem provides a full speed USB1.1 interface to the host as well as an Inter SoPEC Interface (ISI) to other SoPEC devices.

9.2.2 DRAM Subsystem

The DRAM subsystem accepts requests from the CPU, Serial Communications Block (SCB) and blocks within the PEP subsystem. The DRAM subsystem (in particular the DIU) arbitrates the various requests and determines which request should win access to the DRAM. The DIU arbitrates based on configured parameters, to allow sufficient access to DRAM for all requesters. The DIU also hides the implementation specifics of the DRAM such as page size, number of banks, refresh rates etc.

9.2.3 Print Engine Pipeline (PEP) Subsystem

The Print Engine Pipeline (PEP) subsystem accepts compressed pages from DRAM and renders them to bi-level dots for a given print line destined for a printhead interface that communicates directly with up to 2 segments of a bi-lithic printhead.

The first stage of the page expansion pipeline is the CDU, LBD and TE. The CDU expands the JPEG-compressed contone (typically CMYK) layer, the LBD expands the compressed bi-level layer (typically K), and the TE encodes Netpage tags for later rendering (typically in IR or K ink). The output from the first stage is a set of buffers: the CFU, SFU, and TFU. The CFU and SFU buffers are implemented in DRAM.

The second stage is the HCU, which dithers the contone layer, and composites position tags and the bi-level spot0 layer over the resulting bi-level dithered layer. A number of options exist for the way in which compositing occurs. Up to 6 channels of bi-level data are produced from this stage. Note that not all 6 channels may be present on the printhead. For example, the printhead may be CMY only, with K pushed into the CMY channels and IR ignored. Alternatively, the position tags may be printed in K if IR ink is not available (or for testing purposes).

The third stage (DNC) compensates for dead nozzles in the printhead by color redundancy and error diffusing dead nozzle data into surrounding dots.

The resultant bi-level 6 channel dot-data (typically CMYK-IRF) is buffered and written out to a set of line buffers stored in DRAM via the DWU.

Finally, the dot-data is loaded back from DRAM, and passed to the printhead interface via a dot FIFO. The dot FIFO accepts data from the LLU at the system clock rate (pclk), while the PHI removes data from the FIFO and sends it to the printhead at a rate of ⅔ times the system clock rate (see Section 9.1).

9.3 SoPEC Block Description

Looking at FIG. 13, the various units are described here in summary form:

TABLE 9 Units within SoPEC Unit Subsystem Acronym Unit Name Description DRAM DIU DRAM interface unit Provides the interface for DRAM read and write access for the various SoPEC units, CPU and the SCB block. The DIU provides arbitration between competing units controls DRAM access. DRAM Embedded DRAM 20 Mbits of embedded DRAM, CPU CPU Central Processing CPU for system configuration and control Unit MMU Memory Management Limits access to certain memory address areas Unit in CPU user mode RDU Real-time Debug Unit Facilitates the observation of the contents of most of the CPU addressable registers in SoPEC in addition to some pseudo-registers in realtime. TIM General Timer Contains watchdog and general system timers LSS Low Speed Serial Low level controller for interfacing with the QA Interfaces chips GPIO General Purpose IOs General IO controller, with built-in Motor control unit, LED pulse units and de-glitch circuitry ROM Boot ROM 16 KBytes of System Boot ROM code ICU Interrupt Controller General Purpose interrupt controller with Unit configurable priority, and masking. CPR Clock, Power and Central Unit for controlling and generating the Reset block system clocks and resets and powerdown mechanisms PSS Power Save Storage Storage retained while system is powered down USB Universal Serial Bus USB device controller for interfacing with the Device host USB. ISI Inter-SoPEC Interface ISI controller for data and control communication with other SoPEC's in a multi- SoPEC system SCB Serial Communication Contains both the USB and ISI blocks. Block Print Engine PCU PEP controller Provides external CPU with the means to read Pipeline and write PEP Unit registers, and read and (PEP) write DRAM in single 32-bit chunks. CDU Contone decoder unit Expands JPEG compressed contone layer and writes decompressed contone to DRAM CFU Contone FIFO Unit Provides line buffering between CDU and HCU LBD Lossless Bi-level Expands compressed bi-level layer. Decoder SFU Spot FIFO Unit Provides line buffering between LBD and HCU TE Tag encoder Encodes tag data into line of tag dots. TFU Tag FIFO Unit Provides tag data storage between TE and HCU HCU Halftoner compositor Dithers contone layer and composites the bi- unit level spot 0 and position tag dots. DNC Dead Nozzle Compensates for dead nozzles by color Compensator redundancy and error diffusing dead nozzle data into surrounding dots. DWU Dotline Writer Unit Writes out the 6 channels of dot data for a given printline to the line store DRAM LLU Line Loader Unit Reads the expanded page image from line store, formatting the data appropriately for the bi-lithic printhead. PHI PrintHead Interface Is responsible for sending dot data to the bi- lithic printheads and for providing line synchronization between multiple SoPECs. Also provides test interface to printhead such as temperature monitoring and Dead Nozzle Identification. 9.4 Addressing Scheme in SoPEC

SoPEC must address

-   -   20 Mbit DRAM.     -   PCU addressed registers in PEP.     -   CPU-subsystem addressed registers.

SoPEC has a unified address space with the CPU capable of addressing all CPU-subsystem and PCU-bus accessible registers (in PEP) and all locations in DRAM. The CPU generates byte-aligned addresses for the whole of SoPEC.

22 bits are sufficient to byte address the whole SoPEC address space.

9.4.1 DRAM Addressing Scheme

The embedded DRAM is composed of 256-bit words. However the CPU-subsystem may need to write individual bytes of DRAM. Therefore it was decided to make the DIU byte addressable. 22 bits are required to byte address 20 Mbits of DRAM. Most blocks read or write 256-bit words of DRAM. Therefore only the top 17 bits i.e. bits 21 to 5 are required to address 256-bit word aligned locations.

The exceptions are

-   -   CDU which can write 64-bits so only the top 19 address bits i.e.         bits 21-3 are required.     -   The CPU-subsystem always generates a 22-bit byte-aligned DIU         address but it will send flags to the DIU indicating whether it         is an 8, 16 or 32-bit write.

All DIU accesses must be within the same 256-bit aligned DRAM word.

9.4.2 PEP Unit DRAM Addressing

PEP Unit configuration registers which specify DRAM locations should specify 256-bit aligned DRAM addresses i.e. using address bits 21:5. Legacy blocks from PEC1 e.g. the LBD and TE may need to specify 64-bit aligned DRAM addresses if these reused blocks DRAM addressing is difficult to modify. These 64-bit aligned addresses require address bits 21:3. However, these 64-bit aligned addresses should be programmed to start at a 256-bit DRAM word boundary.

Unlike PEC1, there are no constraints in SoPEC on data organization in DRAM except that all data structures must start on a 256-bit DRAM boundary. If data stored is not a multiple of 256-bits then the last word should be padded.

9.4.3 CPU Subsystem Bus Addressed Registers

The CPU subsystem bus supports 32-bit word aligned read and write accesses with variable access timings. See section 11.4 for more details of the access protocol used on this bus. The CPU subsystem bus does not currently support byte reads and writes but this can be added at a later date if required by imported IP.

9.4.4 PCU Addressed Registers in PEP

The PCU only supports 32-bit register reads and writes for the PEP blocks. As the PEP blocks only occupy a subsection of the overall address map and the PCU is explicitly selected by the MMU when a PEP block is being accessed the PCU does not need to perform a decode of the higher-order address bits. See Table 11 for the PEP subsystem address map.

9.5 SoPEC Memory Map

9.5.1 Main Memory Map

The system wide memory map is shown in FIG. 14 below. The memory map is discussed in detail in Section 11 11 Central Processing Unit (CPU).

9.5.2 CPU-Bus Peripherals Address Map

The address mapping for the peripherals attached to the CPU-bus is shown in Table 10 below. The MMU performs the decode of cpu_adr[21:12] to generate the relevant cpu_block_select signal for each block. The addressed blocks decode however many of the lower order bits of cpu_adr[11:2] are required to address all the registers within the block.

TABLE 10 CPU-bus peripherals address map Block_base Address ROM_base 0x0000_0000 MMU_base 0x0001_0000 TIM_base 0x0001_1000 LSS_base 0x0001_2000 GPIO_base 0x0001_3000 SCB_base 0x0001_4000 ICU_base 0x0001_5000 CPR_base 0x0001_6000 DIU_base 0x0001_7000 PSS_base 0x0001_8000 Reserved 0x0001_9000 to 0x0001_FFFF PCU_base 0x0002_0000 to 0x0002_BFFF 9.5.3 PCU Mapped Registers (PEP Blocks) Address Map

The PEP blocks are addressed via the PCU. From FIG. 14, the PCU mapped registers are in the range 0x0002_(—)0000 to 0x0002_BFFF. From Table 11 it can be seen that there are 12 sub-blocks within the PCU address space. Therefore, only four bits are necessary to address each of the sub-blocks within the PEP part of SoPEC. A further 12 bits may be used to address any configurable register within a PEP block. This gives scope for 1024 configurable registers per sub-block (the PCU mapped registers are all 32-bit addressed registers so the upper 10 bits are required to individually address them). This address will come either from the CPU or from a command stored in DRAM. The bus is assembled as follows:

-   -   address[15:12]=sub-block address,     -   address[n:2]=register address within sub-block, only the number         of bits required to decode the registers within each sub-block         are used,     -   address[1:0]=byte address, unused as PCU mapped registers are         all 32-bit addressed registers.

So for the case of the HCU, its addresses range from 0x7000 to 0x7FFF within the PEP subsystem or from 0x0002_(—)7000 to 0x0002_(—)7FFF in the overall system.

TABLE 11 PEP blocks address map Block_base Address PCU_base 0x0002_0000 CDU_base 0x0002_1000 CFU_base 0x0002_2000 LBD_base 0x0002_3000 SFU_base 0x0002_4000 TE_base 0x0002_5000 TFU_base 0x0002_6000 HCU_base 0x0002_7000 DNC_base 0x0002_8000 DWU_base 0x0002_9000 LLU_base 0x0002_A000 PHI_base 0x0002_B000 to 0x0002_BFFF 9.6 Buffer Management in SoPEC

As outlined in Section 9.1, SoPEC has a requirement to print 1 side every 2 seconds i.e. 30 sides per minute.

9.6.1 Page Buffering

Approximately 2 Mbytes of DRAM are reserved for compressed page buffering in SoPEC. If a page is compressed to fit within 2 Mbyte then a complete page can be transferred to DRAM before printing. However, the time to transfer 2 Mbyte using USB 1.1 is approximately 2 seconds. The worst case cycle time to print a page then approaches 4 seconds. This reduces the worst-case print speed to 15 pages per minute.

9.6.2 Band Buffering

The SoPEC page-expansion blocks support the notion of page banding. The page can be divided into bands and another band can be sent down to SoPEC while we are printing the current band. Therefore we can start printing once at least one band has been downloaded.

The band size granularity should be carefully chosen to allow efficient use of the USB bandwidth and DRAM buffer space. It should be small enough to allow seamless 30 sides per minute printing but not so small as to introduce excessive CPU overhead in orchestrating the data transfer and parsing the band headers. Band-finish interrupts have been provided to notify the CPU of free buffer space. It is likely that the host PC will supervise the band transfer and buffer management instead of the SoPEC CPU.

If SoPEC starts printing before the complete page has been transferred to memory there is a risk of a buffer underrun occurring if subsequent bands are not transferred to SoPEC in time e.g. due to insufficient USB bandwidth caused by another USB peripheral consuming USB bandwidth. A buffer underrun occurs if a line synchronisation pulse is received before a line of data has been transferred to the printhead and causes the print job to fail at that line. If there is no risk of buffer underrun then printing can safely start once at least one band has been downloaded.

If there is a risk of a buffer underrun occurring due to an interruption of compressed page data transfer, then the safest approach is to only start printing once we have loaded up the data for a complete page. This means that a worst case latency in the region of 2 seconds (with USB1.1) will be incurred before printing the first page. Subsequent pages will take 2 seconds to print giving us the required sustained printing rate of 30 sides per minute.

A Storage SoPEC (Section 7.2.5) could be added to the system to provide guaranteed bandwidth data delivery. The print system could also be constructed using an ISI-Bridge chip (Section 7.2.6) to provide guaranteed data delivery.

The most efficient page banding strategy is likely to be determined on a per page/print job basis and so SoPEC will support the use of bands of any size.

10 SoPEC Use Cases

10.1 Introduction

This chapter is intended to give an overview of a representative set of scenarIOs or use cases which SoPEC can perform. SoPEC is by no means restricted to the particular use cases described and not every SoPEC system is considered here.

In this chapter we discuss SoPEC use cases under four headings:

-   1) Normal operation use cases. -   2) Security use cases. -   3) Miscellaneous use cases. -   4) Failure mode use cases.

Use cases for both single and multi-SoPEC systems are outlined.

Some tasks may be composed of a number of sub-tasks.

The realtime requirements for SoPEC software tasks are discussed in “11 Central Processing Unit (CPU)” under Section 11.3 Realtime requirements.

10.2 Normal Operation in a single SoPEC System with USB Host Connection

SoPEC operation is broken up into a number of sections which are outlined below. Buffer management in a SoPEC system is normally performed by the host.

10.2.1 Powerup

Powerup describes SoPEC initialisation following an external reset or the watchdog timer system reset.

A typical powerup sequence is:

-   1) Execute reset sequence for complete SoPEC. -   2) CPU boot from ROM. -   3) Basic configuration of CPU peripherals, SCB and DIU. DRAM     initialisation. USB Wakeup. -   4) Download and authentication of program (see Section 10.5.2). -   5) Execution of program from DRAM. -   6) Retrieve operating parameters from PRINTER_QA and authenticate     operating parameters. -   7) Download and authenticate any further datasets.     10.2.2 USB Wakeup

The CPU can put different sections of SoPEC into sleep mode by writing to registers in the CPR block (chapter 16). Normally the CPU sub-system and the DRAM will be put in sleep mode but the SCB and power-safe storage (PSS) will still be enabled.

Wakeup describes SoPEC recovery from sleep mode with the SCB and power-safe storage (PSS) still enabled. In a single SoPEC system, wakeup can be initiated following a USB reset from the SCB.

A typical USB wakeup sequence is:

-   1) Execute reset sequence for sections of SoPEC in sleep mode. -   2) CPU boot from ROM, if CPU-subsystem was in sleep mode. -   3) Basic configuration of CPU peripherals and DIU, and DRAM     initialisation, if required. -   4) Download and authentication of program using results in     Power-Safe Storage (PSS) (see Section 10.5.2). -   5) Execution of program from DRAM. -   6) Retrieve operating parameters from PRINTER_QA and authenticate     operating parameters. -   7) Download and authenticate using results in PSS of any further     datasets (programs).     10.2.3 Print Initialization

This sequence is typically performed at the start of a print job following powerup or wakeup:

-   1) Check amount of ink remaining via QA chips. -   2) Download static data e.g. dither matrices, dead nozzle tables     from host to DRAM. -   3) Check printhead temperature, if required, and configure printhead     with firing pulse profile etc. accordingly. -   4) Initiate printhead pre-heat sequence, if required.     10.2.4 First Page Download

Buffer management in a SoPEC system is normally performed by the host.

First page, first band download and processing:

-   1) The host communicates to the SoPEC CPU over the USB to check that     DRAM space remaining is sufficient to download the first band. -   2) The host downloads the first band (with the page header) to DRAM. -   3) When the complete page header has been downloaded the SoPEC CPU     processes the page header, calculates PEP register commands and     writes directly to PEP registers or to DRAM. -   4) If PEP register commands have been written to DRAM, execute PEP     commands from DRAM via PCU.

Remaining bands download and processing:

-   1) Check DRAM space remaining is sufficient to download the next     band. -   2) Download the next band with the band header to DRAM. -   3) When the complete band header has been downloaded, process the     band header according to whichever band-related register updating     mechanism is being used.     10.2.5 Start Printing -   1) Wait until at least one band of the first page has been     downloaded. One approach is to only start printing once we have     loaded up the data for a complete page. If we start printing before     the complete page has been transferred to memory we run the risk of     a buffer underrun occurring because compressed page data was not     transferred to SoPEC in time e.g. due to insufficient USB bandwidth     caused by another USB peripheral consuming USB bandwidth. -   2) Start all the PEP Units by writing to their Go registers, via PCU     commands executed from DRAM or direct CPU writes. A rapid startup     order for the PEP units is outlined in Table 12.

TABLE 12 Typical PEP Unit startup order for printing a page. Step# Unit 1 DNC 2 DWU 3 HCU 4 PHI 5 LLU 6 CFU, SFU, TFU 7 CDU 8 TE, LBD

-   3) Print ready interrupt occurs (from PHI). -   4) Start motor control, if first page, otherwise feed the next page.     This step could occur before the print ready interrupt. -   5) Drive LEDs, monitor paper status. -   6) Wait for page alignment via page sensor(s) GPIO interrupt. -   7) CPU instructs PHI to start producing line syncs and hence     commence printing, or wait for an external device to produce line     syncs. -   8) Continue to download bands and process page and band headers for     next page.     10.2.6 Next Page(s) Download

As for first page download, performed during printing of current page.

10.2.7 Between Bands

When the finished band flags are asserted band related registers in the CDU, LBD, TE need to be re-programmed before the subsequent band can be printed. This can be via PCU commands from DRAM. Typically only 3-5 commands per decompression unit need to be executed. These registers can also be reprogrammed directly by the CPU or most likely by updating from shadow registers. The finished band flag interrupts the CPU to tell the CPU that the area of memory associated with the band is now free.

10.2.8 During Page Print

Typically during page printing ink usage is communicated to the QA chips.

-   1) Calculate ink printed (from PHI). -   2) Decrement ink remaining (via QA chips). -   3) Check amount of ink remaining (via QA chips). This operation may     be better performed while the page is being printed rather than at     the end of the page.     10.2.9 Page Finish

These operations are typically performed when the page is finished:

-   1) Page finished interrupt occurs from PHI. -   2) Shutdown the PEP blocks by de-asserting their Go registers. A     typical shutdown order is defined in Table 13. This will set the PEP     Unit state-machines to their idle states without resetting their     configuration registers. -   3) Communicate ink usage to QA chips, if required.

TABLE 13 End of page shutdown order for PEP Units. Step# Unit 1 PHI (will shutdown by itself in the normal case at the end of a page) 2 DWU (shutting this down stalls the DNC and therefore the HCU and above) 3 LLU (should already be halted due to PHI at end of last line of page) 4 TE (this is the only dot supplier likely to be running, halted by the HCU) 5 CDU (this is likely to already be halted due to end of contone band) 6 CPU, SFU, TFU, LBD (order unimportant, and should already be halted due to end of band) 7 HCU, DNC (order unimportant, should already have halted) 10.2.10 Start of Next Page

These operations are typically performed before printing the next page:

-   1) Re-program the PEP Units via PCU command processing from DRAM     based on page header. -   2) Go to Start printing.     10.2.11 End of Document -   1) Stop motor control.     10.2.12 Sleep Mode

The CPU can put different sections of SoPEC into sleep mode by writing to registers in the CPR block described in Section 16.

-   1) Instruct host PC via USB that SoPEC is about to sleep. -   2) Store reusable authentication results in Power-Safe Storage     (PSS). -   3) Put SoPEC into defined sleep mode.     10.3 Normal Operation in a Multi-SoPEC System—ISIMaster SoPEC

In a multi-SoPEC system the host generally manages program and compressed page download to all the SoPECs. Inter-SoPEC communication is over the ISI link which will add a latency. In the case of a multi-SoPEC system with just one USB 1.1 connection, the SoPEC with the USB connection is the ISIMaster. The ISI-bridge chip is the ISIMaster in the case of an ISI-Bridge SoPEC configuration. While it is perfectly possible for an ISISlave to have a direct USB connection to the host we do not treat this scenario explicitly here to avoid possible confusion.

In a multi-SoPEC system one of the SoPECs will be the PrintMaster. This SoPEC must manage and control sensors and actuators e.g. motor control. These sensors and actuators could be distributed over all the SoPECs in the system. An ISIMaster SoPEC may also be the PrintMaster SoPEC.

In a multi-SoPEC system each printing SoPEC will generally have its own PRINTER_QA chip (or at least access to a PRINTER_QA chip that contains the SoPEC's SoPEC_id_key) to validate operating parameters and ink usage. The results of these operations may be communicated to the PrintMaster SoPEC.

In general the ISIMaster may need to be able to:

Send messages to the ISISlaves which will cause the ISISlaves to send their status to the ISIMaster.

Instruct the ISISlaves to perform certain operations.

As the ISI is an insecure interface commands issued over the ISI are regarded as user mode commands. Supervisor mode code running on the SoPEC CPUs will allow or disallow these commands. The software protocol needs to be constructed with this in mind.

The ISIMaster will initiate all communication with the ISISlaves.

SoPEC operation is broken up into a number of sections which are outlined below.

10.3.1 Powerup

Powerup describes SoPEC initialisation following an external reset or the watchdog timer system reset.

-   1) Execute reset sequence for complete SoPEC. -   2) CPU boot from ROM. -   3) Basic configuration of CPU peripherals, SCB and DIU. DRAM     initialisation USB Wakeup -   4) SoPEC identification by activity on USB end-points 2-4 indicates     it is the ISIMaster (unless the SoPEC CPU has explicitly disabled     this function). -   5) Download and authentication of program (see Section 10.5.3). -   6) Execution of program from DRAM. -   7) Retrieve operating parameters from PRINTER_QA and authenticate     operating parameters. -   8) Download and authenticate any further datasets (programs). -   9) The initial dataset may be broadcast to all the ISISlaves. -   10) ISIMaster master SoPEC then waits for a short time to allow the     authentication to take place on the ISISlave SoPECs. -   11) Each ISISlave SoPEC is polled for the result of its program code     authentication process. -   12) If all ISISlaves report successful authentication the OEM code     module can be distributed and authenticated. OEM code will most     likely reside on one SoPEC.     10.3.2 USB Wakeup

The CPU can put different sections of SoPEC into sleep mode by writing to registers in the CPR block [16]. Normally the CPU sub-system and the DRAM will be put in sleep mode but the SCB and power-safe storage (PSS) will still be enabled.

Wakeup describes SoPEC recovery from sleep mode with the SCB and power-safe storage (PSS) still enabled. For an ISIMaster SoPEC connected to the host via USB, wakeup can be initiated following a USB reset from the SCB.

A typical USB wakeup sequence is:

-   1) Execute reset sequence for sections of SoPEC in sleep mode. -   2) CPU boot from ROM, if CPU-subsystem was in sleep mode. -   3) Basic configuration of CPU peripherals and DIU, and DRAM     initialisation, if required. -   4) SoPEC identification by activity on USB end-points 2-4 indicates     it is the ISIMaster (unless the SoPEC CPU has explicitly disabled     this function). -   5) Download and authentication of program using results in     Power-Safe Storage (PSS) (see Section 10.5.3). -   6) Execution of program from DRAM. -   7) Retrieve operating parameters from PRINTER_QA and authenticate     operating parameters. -   8) Download and authenticate any further datasets (programs) using     results in Power-Safe Storage (PSS) (see Section 10.5.3). -   9) Following steps as per Powerup.     10.3.3 Print Initialization

This sequence is typically performed at the start of a print job following powerup or wakeup:

-   1) Check amount of ink remaining via QA chips which may be present     on a ISISlave SoPEC. -   2) Download static data e.g. dither matrices, dead nozzle tables     from host to DRAM. -   3) Check printhead temperature, if required, and configure printhead     with firing pulse profile etc. accordingly. Instruct ISISlaves to     also perform this operation. -   4) Initiate printhead pre-heat sequence, if required. Instruct     ISISlaves to also perform this operation     10.3.4 First Page Download

Buffer management in a SoPEC system is normally performed by the host.

-   1) The host communicates to the SoPEC CPU over the USB to check that     DRAM space remaining is sufficient to download the first band. -   2) The host downloads the first band (with the page header) to DRAM. -   3) When the complete page header has been downloaded the SoPEC CPU     processes the page header, calculates PEP register commands and     write directly to PEP registers or to DRAM. -   4) If PEP register commands have been written to DRAM, execute PEP     commands from DRAM via PCU.

Poll ISISlaves for DRAM status and download compressed data to ISISlaves.

Remaining first page bands download and processing:

-   1) Check DRAM space remaining is sufficient to download the next     band. -   2) Download the next band with the band header to DRAM. -   3) When the complete band header has been downloaded, process the     band header according to whichever band-related register updating     mechanism is being used.

Poll ISISlaves for DRAM status and download compressed data to ISISlaves.

10.3.5 Start Printing

-   1) Wait until at least one band of the first page has been     downloaded. -   2) Start all the PEP Units by writing to their Go registers, via PCU     commands executed from DRAM or direct CPU writes, in the suggested     order defined in Table. -   3) Print ready interrupt occurs (from PHI). Poll ISISlaves until     print ready interrupt. -   4) Start motor control (which may be on an ISISlave SoPEC), if first     page, otherwise feed the next page. This step could occur before the     print ready interrupt. -   5) Drive LEDS, monitor paper status (which may be on an ISISlave     SoPEC). -   6) Wait for page alignment via page sensor(s) GPIO interrupt (which     may be on an ISISlave SoPEC). -   7) If the LineSyncMaster is a SoPEC its CPU instructs PHI to start     producing master line syncs. Otherwise wait for an external device     to produce line syncs. -   8) Continue to download bands and process page and band headers for     next page.     10.3.6 Next Page(s) Download

As for first page download, performed during printing of current page.

10.3.7 Between Bands

When the finished band flags are asserted band related registers in the CDU, LBD and TE need to be re-programmed. This can be via PCU commands from DRAM. Typically only 3-5 commands per decompression unit need to be executed. These registers can also be reprogrammed directly by the CPU or by updating from shadow registers. The finished band flag interrupts to the CPU, tell the CPU that the area of memory associated with the band is now free.

10.3.8 During Page Print

Typically during page printing ink usage is communicated to the QA chips.

-   1) Calculate ink printed (from PHI). -   2) Decrement ink remaining (via QA chips). -   3) Check amount of ink remaining (via QA chips). This operation may     be better performed while the page is being printed rather than at     the end of the page.     10.3.9 Page Finish

These operations are typically performed when the page is finished:

-   1) Page finished interrupt occurs from PHI. Poll ISISlaves for page     finished interrupts. -   2) Shutdown the PEP blocks by de-asserting their Go registers in the     suggested order in Table. This will set the PEP Unit state-machines     to their startup states. -   3) Communicate ink usage to QA chips, if required.     10.3.10 Start of Next Page

These operations are typically performed before printing the next page:

-   1) Re-program the PEP Units via PCU command processing from DRAM     based on page header. -   2) Go to Start printing.     10.3.11 End of Document -   1) Stop motor control. This may be on an ISISlave SoPEC.     10.3.12 Sleep Mode

The CPU can put different sections of SoPEC into sleep mode by writing to registers in the CPR block [16]. This may be as a result of a command from the host or as a result of a timeout.

-   1) Inform host PC of which parts of SoPEC system are about to sleep. -   2) Instruct ISISlaves to enter sleep mode. -   3) Store reusable cryptographic results in Power-Safe Storage (PSS). -   4) Put ISIMaster SoPEC into defined sleep mode.     10.4 Normal Operation in a Multi-SoPEC System—ISISlave SoPEC

This section the outline typical operation of an ISISlave SoPEC in a multi-SoPEC system. The ISIMaster can be another SoPEC or an ISI-Bridge chip. The ISISlave communicates with the host either via the ISIMaster or using a direct connection such as USB. For this use case we consider only an ISISlave that does not have a direct host connection. Buffer management in a SoPEC system is normally performed by the host.

10.4.1 Powerup

Powerup describes SoPEC initialisation following an external reset or the watchdog timer system reset.

A typical powerup sequence is:

-   1) Execute reset sequence for complete SoPEC. -   2) CPU boot from ROM. -   3) Basic configuration of CPU peripherals, SCB and DIU. DRAM     initialisation. -   4) Download and authentication of program (see Section 10.5.3). -   5) Execution of program from DRAM. -   6) Retrieve operating parameters from PRINTER_QA and authenticate     operating parameters. -   7) SoPEC identification by sampling GPIO pins to determine ISIId.     Communicate ISIId to ISIMaster. -   8) Download and authenticate any further datasets.     10.4.2 ISI Wakeup

The CPU can put different sections of SoPEC into sleep mode by writing to registers in the CPR block [16]. Normally the CPU sub-system and the DRAM will be put in sleep mode but the SCB and power-safe storage (PSS) will still be enabled.

Wakeup describes SoPEC recovery from sleep mode with the SCB and power-safe storage (PSS) still enabled. In an ISISlave SoPEC, wakeup can be initiated following an ISI reset from the SCB. A typical ISI wakeup sequence is:

-   1) Execute reset sequence for sections of SoPEC in sleep mode. -   2) CPU boot from ROM, if CPU-subsystem was in sleep mode. -   3) Basic configuration of CPU peripherals and DIU, and DRAM     initialisation, if required. -   4) Download and authentication of program using results in     Power-Safe Storage (PSS) (see Section 10.5.3). -   5) Execution of program from DRAM. -   6) Retrieve operating parameters from PRINTER_QA and authenticate     operating parameters. -   7) SoPEC identification by sampling GPIO pins to determine ISIId.     Communicate ISIId to ISIMaster. -   8) Download and authenticate any further datasets.     10.4.3 Print Initialization

This sequence is typically performed at the start of a print job following powerup or wakeup:

-   1) Check amount of ink remaining via QA chips. -   2) Download static data e.g. dither matrices, dead nozzle tables     from ISI to DRAM. -   3) Check printhead temperature, if required, and configure printhead     with firing pulse profile etc. accordingly. -   4) Initiate printhead pre-heat sequence, if required.     10.4.4 First Page Download

Buffer management in a SoPEC system is normally performed by the host via the ISI.

-   1) Check DRAM space remaining is sufficient to download the first     band. -   2) The host downloads the first band (with the page header) to DRAM     via the ISI. -   3) When the complete page header has been downloaded, process the     page header, calculate PEP register commands and write directly to     PEP registers or to DRAM. -   4) If PEP register commands have been written to DRAM, execute PEP     commands from DRAM via PCU.

Remaining first page bands download and processing:

-   1) Check DRAM space remaining is sufficient to download the next     band. -   2) The host downloads the first band (with the page header) to DRAM     via the ISI. -   3) When the complete band header has been downloaded, process the     band header according to whichever band-related register updating     mechanism is being used.     10.4.5 Start Printing -   1) Wait until at least one band of the first page has been     downloaded. -   2) Start all the PEP Units by writing to their Go registers, via PCU     commands executed from DRAM or direct CPU writes, in the order     defined in Table -   3) Print ready interrupt occurs (from PHI). Communicate to     PrintMaster via ISI. -   4) Start motor control, if attached to this ISISlave, when requested     by PrintMaster, if first page, otherwise feed next page. This step     could occur before the print ready interrupt -   5) Drive LEDS, monitor paper status, if on this ISISlave SoPEC, when     requested by PrintMaster -   6) Wait for page alignment via page sensor(s) GPIO interrupt, if on     this ISISlave SoPEC, and send to PrintMaster. -   7) Wait for line sync and commence printing. -   8) Continue to download bands and process page and band headers for     next page.     10.4.6 Next Page(s) Download

As for first band download, performed during printing of current page.

10.4.7 Between Bands

When the finished band flags are asserted band related registers in the CDU, LBD and TE need to be re-programmed. This can be via PCU commands from DRAM. Typically only 3-5 commands per decompression unit need to be executed. These registers can also be reprogrammed directly by the CPU or by updating from shadow registers. The finished band flag interrupts to the CPU tell the CPU that the area of memory associated with the band is now free.

10.4.8 During Page Print

Typically during page printing ink usage is communicated to the QA chips.

-   1) Calculate ink printed (from PHI). -   2) Decrement ink remaining (via QA chips). -   3) Check amount of ink remaining (via QA chips). This operation may     be better performed while the page is being printed rather than at     the end of the page.     10.4.9 Page Finish

These operations are typically performed when the page is finished:

-   1) Page finished interrupt occurs from PHI. Communicate page     finished interrupt to PrintMaster. -   2) Shutdown the PEP blocks by de-asserting their Go registers in the     suggested order in Table. This will set the PEP Unit state-machines     to their startup states. -   3) Communicate ink usage to QA chips, if required.     10.4.10 Start of Next Page

These operations are typically performed before printing the next page:

-   1) Re-program the PEP Units via PCU command processing from DRAM     based on page header. -   2) Go to Start printing.     10.4.11 End of Document

Stop motor control, if attached to this ISISlave, when requested by PrintMaster.

10.4.12 Powerdown

In this mode SoPEC is no longer powered.

-   1) Powerdown ISISlave SoPEC when instructed by ISIMaster.     10.4.13 Sleep

The CPU can put different sections of SoPEC into sleep mode by writing to registers in the CPR block [16]. This may be as a result of a command from the host or ISIMaster or as a result of a timeout.

-   1) Store reusable cryptographic results in Power-Safe Storage (PSS). -   2) Put SoPEC into defined sleep mode.     10.5 Security Use Cases

Please see the ‘SoPEC Security Overview’ [9] document for a more complete description of SoPEC security issues. The SoPEC boot operation is described in the ROM chapter of the SoPEC hardware design specification, Section 17.2.

10.5.1 Communication with the QA Chips

Communication between SoPEC and the QA chips (i.e. INK_QA and PRINTER_QA) will take place on at least a per power cycle and per page basis. Communication with the QA chips has three principal purposes: validating the presence of genuine QA chips (i.e the printer is using approved consumables), validation of the amount of ink remaining in the cartridge and authenticating the operating parameters for the printer. After each page has been printed, SoPEC is expected to communicate the number of dots fired per ink plane to the QA chipset. SoPEC may also initiate decoy communications with the QA chips from time to time.

Process:

-   -   When validating ink consumption SoPEC is expected to principally         act as a conduit between the PRINTER_QA and INK_QA chips and to         take certain actions (basically enable or disable printing and         report status to host PC) based on the result. The communication         channels are insecure but all traffic is signed to guarantee         authenticity.         Known Weaknesses     -   All communication to the QA chips is over the LSS interfaces         using a serial communication protocol. This is open to         observation and so the communication protocol could be reverse         engineered. In this case both the PRINTER_QA and INK_QA chips         could be replaced by impostor devices (e.g. a single FPGA) that         successfully emulated the communication protocol. As this would         require physical modification of each printer this is considered         to be an acceptably low risk. Any messages that are not signed         by one of the symmetric keys (such as the SoPEC_id_key) could be         reverse engineered. The imposter device must also have access to         the appropriate keys to crack the system.     -   If the secret keys in the QA chips are exposed or cracked then         the system, or parts of it, is compromised.         Assumptions:

-   [1] The QA chips are not involved in the authentication of     downloaded SoPEC code

-   [2] The QA chip in the ink cartridge (INK_QA) does not directly     affect the operation of the cartridge in any way i.e. it does not     inhibit the flow of ink etc.

-   [3] The INK_QA and PRINTER_QA chips are identical in their virgin     state. They only become a INK_QA or PRINTER_QA after their FlashROM     has been programmed.     10.5.2 Authentication of Downloaded Code in a Single SoPEC System     Process:

-   1) SoPEC identification by activity on USB end-points 2-4 indicates     it is the ISIMaster (unless the SoPEC CPU has explicitly disabled     this function).

-   2) The program is downloaded to the embedded DRAM.

-   3) The CPU calculates a SHA-1 hash digest of the downloaded program.

-   4) The ResetSrc register in the CPR block is read to determine     whether or not a power-on reset occurred.

-   5) If a power-on reset occurred the signature of the downloaded code     (which needs to be in a known location such as the first or last N     bytes of the downloaded code) is decrypted using the Silverbrook     public boot0key stored in ROM. This decrypted signature is the     expected SHA-1 hash of the accompanying program. The encryption     algorithm is likely to be a public key algorithm such as RSA. If a     power-on reset did not occur then the expected SHA-1 hash is     retrieved from the PSS and the compute intensive decryption is not     required.

-   6) The calculated and expected hash values are compared and if they     match then the programs authenticity has been verified.

-   7) If the hash values do not match then the host PC is notified of     the failure and the SoPEC will await a new program download.

-   8) If the hash values match then the CPU starts executing the     downloaded program.

-   9) If, as is very likely, the downloaded program wishes to download     subsequent programs (such as OEM code) it is responsible for     ensuring the authenticity of everything it downloads. The downloaded     program may contain public keys that are used to authenticate     subsequent downloads, thus forming a hierarchy of authentication.     The SoPEC ROM does not control these authentications—it is solely     concerned with verifying that the first program downloaded has come     from a trusted source.

-   10) At some subsequent point OEM code starts executing. The     Silverbrook supervisor code acts as an O/S to the OEM user mode     code. The OEM code must access most SoPEC functionality via system     calls to the Silverbrook code.

-   11) The OEM code is expected to perform some simple ‘turn on the     lights’ tasks after which the host PC is informed that the printer     is ready to print and the Start Printing use case comes into play.     Known Weaknesses:     -   If the Silverbrook private boot0key is exposed or cracked then         the system is seriously compromised. A ROM mask change would be         required to reprogram the boot0key.         10.5.3 Authentication of Downloaded Code in a Multi-SoPEC System         10.5.3.1 ISIMaster SoPEC Process:

-   1) SoPEC identification by activity on USB end-points 2-4 indicates     it is the ISIMaster.

-   2) The SCB is configured to broadcast the data received from the     host PC.

-   3) The program is downloaded to the embedded DRAM and broadcasted to     all ISISlave SoPECs over the ISI.

-   4) The CPU calculates a SHA-1 hash digest of the downloaded program.

-   5) The ResetSrc register in the CPR block is read to determine     whether or not a power-on reset occurred.

-   6) If a power-on reset occurred the signature of the downloaded code     (which needs to be in a known location such as the first or last N     bytes of the downloaded code) is decrypted using the Silverbrook     public boot0key stored in ROM. This decrypted signature is the     expected SHA-1 hash of the accompanying program. The encryption     algorithm is likely to be a public key algorithm such as RSA. If a     power-on reset did not occur then the expected SHA-1 hash is     retrieved from the PSS and the compute intensive decryption is not     required.

-   7) The calculated and expected hash values are compared and if they     match then the programs authenticity has been verified.

-   8) If the hash values do not match then the host PC is notified of     the failure and the SoPEC will await a new program download.

-   9) If the hash values match then the CPU starts executing the     downloaded program.

-   10) It is likely that the downloaded program will poll each ISISlave     SoPEC for the result of its authentication process and to determine     the number of slaves present and their ISIIds.

-   11) If any ISISlave SoPEC reports a failed authentication then the     ISIMaster communicates this to the host PC and the SoPEC will await     a new program download.

-   12) If all ISISlaves report successful authentication then the     downloaded program is responsible for the downloading,     authentication and distribution of subsequent programs within the     multi-SoPEC system.

-   13) At some subsequent point OEM code starts executing. The     Silverbrook supervisor code acts as an O/S to the OEM user mode     code. The OEM code must access most SoPEC functionality via system     calls to the Silverbrook code.

-   14) The OEM code is expected to perform some simple ‘turn on the     lights’ tasks after which the master SoPEC determines that all     SoPECs are ready to print. The host PC is informed that the printer     is ready to print and the Start Printing use case comes into play.     10.5.3.2 ISISlave SoPEC Process:

-   1) When the CPU comes out of reset the SCB will be in slave mode,     and the SCB is already configured to receive data from both the ISI     and USB.

-   2) The program is downloaded (via ISI or USB) to embedded DRAM.

-   3) The CPU calculates a SHA-1 hash digest of the downloaded program.

-   4) The ResetSrc register in the CPR block is read to determine     whether or not a power-on reset occurred.

-   5) If a power-on reset occurred the signature of the downloaded code     (which needs to be in a known location such as the first or last N     bytes of the downloaded code) is decrypted using the Silverbrook     public boot0key stored in ROM. This decrypted signature is the     expected SHA-1 hash of the accompanying program. The encryption     algorithm is likely to be a public key algorithm such as RSA. If a     power-on reset did not occur then the expected SHA-1 hash is     retrieved from the PSS and the compute intensive decryption is not     required.

-   6) The calculated and expected hash values are compared and if they     match then the programs authenticity has been verified.

-   7) If the hash values do not match, then the ISISlave device will     await a new program again

-   8) If the hash values match then the CPU starts executing the     downloaded program.

-   9) It is likely that the downloaded program will communicate the     result of its authentication process to the ISIMaster. The     downloaded program is responsible for determining the SoPECs ISIID,     receiving and authenticating any subsequent programs.

-   10) At some subsequent point OEM code starts executing. The     Silverbrook supervisor code acts as an O/S to the OEM user mode     code. The OEM code must access most SoPEC functionality via system     calls to the Silverbrook code.

-   11) The OEM code is expected to perform some simple ‘turn on the     lights’ tasks after which the master SoPEC is informed that this     slave is ready to print. The Start Printing use case then comes into     play.     Known Weaknesses     -   If the Silverbrook private boot0key is exposed or cracked then         the system is seriously compromised.     -   ISI is an open interface i.e. messages sent over the ISI are in         the clear. The communication channels are insecure but all         traffic is signed to guarantee authenticity. As all         communication over the ISI is controlled by Supervisor code on         both the ISIMaster and ISISlave then this also provides some         protection against software attacks.         10.5.4 Authentication and Upgrade of Operating Parameters for a         Printer

The SoPEC IC will be used in a range of printers with different capabilities (e.g. A3/A4 printing, printing speed, resolution etc.). It is expected that some printers will also have a software upgrade capability which would allow a user to purchase a license that enables an upgrade in their printer's capabilities (such as print speed). To facilitate this it must be possible to securely store the operating parameters in the PRINTER_QA chip, to securely communicate these parameters to the SoPEC and to securely reprogram the parameters in the event of an upgrade. Note that each printing SoPEC (as opposed to a SoPEC that is only used for the storage of data) will have its own PRINTER_QA chip (or at least access to a PRINTER_QA that contains the SoPEC's SoPEC_id_key). Therefore both ISIMaster and ISISlave SoPECs will need to authenticate operating parameters.

Process:

-   1) Program code is downloaded and authenticated as described in     sections 10.5.2 and 10.5.3 above. -   2) The program code has a function to create the SoPEC_id_key from     the unique SoPEC_id that was programmed when the SoPEC was     manufactured. -   3) The SoPEC retrieves the signed operating parameters from its     PRINTER_QA chip. The PRINTER_QA chip uses the SoPEC_id_key (which is     stored as part of the pairing process executed during printhead     assembly manufacture & test) to sign the operating parameters which     are appended with a random number to thwart replay attacks. -   4) The SoPEC checks the signature of the operating parameters using     its SoPEC_id_-key. If this signature authentication process is     successful then the operating parameters are considered valid and     the overall boot process continues. If not the error is reported to     the host PC. -   5) Operating parameters may also be set or upgraded using a second     key, the PrintEngineLicense_key, which is stored on the PRINTER_QA     and used to authenticate the change in operating parameters.     Known Weaknesses:     -   It may be possible to retrieve the unique SoPEC_id by placing         the SoPEC in test mode and scanning it out. It is certainly         possible to obtain it by reverse engineering the device. Either         way the SoPEC_id (and by extension the SoPEC_id_key) so obtained         is valid only for that specific SoPEC and so printers may only         be compromised one at a time by parties with the appropriate         specialised equipment. Furthermore even if the SoPEC_id is         compromised, the other keys in the system, which protect the         authentication of consumables and of program code, are         unaffected.         10.6 Miscellaneous Use Cases

There are many miscellaneous use cases such as the following examples. Software running on the SoPEC CPU or host will decide on what actions to take in these scenarios.

10.6.1 Disconnect/Re-Connect of QA Chips.

-   1) Disconnect of a QA chip between documents or if ink runs out     mid-document. -   2) Re-connect of a QA chip once authenticated e.g. ink cartridge     replacement should allow the system to resume and print the next     document     10.6.2 Page Arrives Before Print Ready Interrupt. -   1) Engage clutch to stop paper until print ready interrupt occurs.     10.6.3 Dead-Nozzle Table Upgrade

This sequence is typically performed when dead nozzle information needs to be updated by performing a printhead dead nozzle test.

-   1) Run printhead nozzle test sequence -   2) Either host or SoPEC CPU converts dead nozzle information into     dead nozzle table. -   3) Store dead nozzle table on host. -   4) Write dead nozzle table to SoPEC DRAM.     10.7 Failure Mode Use Cases     10.7.1 System Errors and Security Violations

System errors and security violations are reported to the SoPEC CPU and host. Software running on the SoPEC CPU or host will then decide what actions to take.

Silverbrook code authentication failure.

-   1) Notify host PC of authentication failure. -   2) Abort print run.

OEM code authentication failure.

-   1) Notify host PC of authentication failure. -   2) Abort print run.

Invalid QA chip(s).

-   1) Report to host PC. -   2) Abort print run.

MMU security violation interrupt.

-   1) This is handled by exception handler. -   2) Report to host PC -   3) Abort print run.

Invalid address interrupt from PCU.

-   1) This is handled by exception handler. -   2) Report to host PC. -   3) Abort print run.

Watchdog timer interrupt.

-   1) This is handled by exception handler. -   2) Report to host PC. -   3) Abort print run.

Host PC does not acknowledge message that SoPEC is about to power down.

-   1) Power down anyway.     10.7.2 Printing Errors

Printing errors are reported to the SoPEC CPU and host. Software running on the host or SoPEC CPU will then decide what actions to take.

Insufficient space available in SoPEC compressed band-store to download a band.

-   1) Report to the host PC.

Insufficient ink to print.

-   1) Report to host PC.

Page not downloaded in time while printing.

-   1) Buffer underrun interrupt will occur. -   2) Report to host PC and abort print run.

JPEG decoder error interrupt.

-   1) Report to host PC.     CPU Subsystem     11 Central Processing Unit (CPU)     11.1 Overview

The CPU block consists of the CPU core, MMU, cache and associated logic. The principal tasks for the program running on the CPU to fulfill in the system are:

Communications:

-   -   Control the flow of data from the USB interface to the DRAM and         ISI     -   Communication with the host via USB or ISI     -   Running the USB device driver         PEP Subsystem Control:     -   Page and band header processing (may possibly be performed on         host PC)     -   Configure printing options on a per band, per page, per job or         per power cycle basis     -   Initiate page printing operation in the PEP subsystem     -   Retrieve dead nozzle information from the printhead interface         (PHI) and forward to the host PC     -   Select the appropriate firing pulse profile from a set of         predefined profiles based on the printhead characteristics     -   Retrieve printhead temperature via the PHI         Security:     -   Authenticate downloaded program code     -   Authenticate printer operating parameters     -   Authenticate consumables via the PRINTER_QA and INK_QA chips     -   Monitor ink usage     -   Isolation of OEM code from direct access to the system resources         Other:     -   Drive the printer motors using the GPIO pins     -   Monitoring the status of the printer (paper jam, tray empty         etc.)     -   Driving front panel LEDs     -   Perform post-boot initialisation of the SoPEC device     -   Memory management (likely to be in conjunction with the host PC)     -   Miscellaneous housekeeping tasks

To control the Print Engine Pipeline the CPU is required to provide a level of performance at least equivalent to a 16-bit Hitachi H8-3664 microcontroller running at 16 MHz. An as yet undetermined amount of additional CPU performance is needed to perform the other tasks, as well as to provide the potential for such activity as Netpage page assembly and processing, RIPing etc. The extra performance required is dominated by the signature verification task and the SCB (including the USB) management task. An operating system is not required at present. A number of CPU cores have been evaluated and the LEON P1754 is considered to be the most appropriate solution. A diagram of the CPU block is shown in FIG. 15 below.

11.2 Definitions of I/Os

TABLE 14 CPU Subsystem I/Os Port name Pins I/O Description Clocks and Resets prst_n 1 In Global reset. Synchronous to pclk, active low. Pclk 1 In Global clock CPU to DIU DRAM interface cpu_adr[21:2] 20 Out Address bus for both DRAM and peripheral access cpu_dataout[31:0] 32 Out Data out to both DRAM and peripheral devices. This should be driven at the same time as the cpu_adr and request signals. dram_cpu_data[255:0] 256 In Read data from the DRAM cpu_diu_rreq 1 Out Read request to the DIU DRAM diu_cpu_rack 1 In Acknowledge from DIU that read request has been accepted. diu_cpu_rvalid 1 In Signal from DIU telling SoPEC Unit that valid read data is on the dram_cpu_data bus cpu_diu_wdatavalid 1 Out Signal from the CPU to the DIU indicating that the data currently on the cpu_diu_wdata bus is valid and should be committed to the DIU posted write buffer diu_cpu_write_rdy 1 In Signal from the DIU indicating that the posted write buffer is empty cpu_diu_wdadr[21:4] 18 Out Write address bus to the DIU cpu_diu_wdata[127:0] 128 Out Write data bus to the DIU cpu_diu_wmask[15:0] 16 Out Write mask for the cpu_diu_wdata bus. Each bit corresponds to a byte of the 128-bit cpu_diu_wdata bus. CPU to peripheral blocks cpu_rwn 1 Out Common read/not-write signal from the CPU cpu_acode[1:0] 2 Out CPU access code signals. cpu_acode[0] - Program (0)/Data (1) access cpu_acode[1] - User (0)/Supervisor (1) access cpu_cpr_sel 1 Out CPR block select. cpr_cpu_rdy 1 In Ready signal to the CPU. When cpr_cpu_rdy is high it indicates the last cycle of the access. For a write cycle this means cpu_dataout has been registered by the CPR block and for a read cycle this means the data on cpr_cpu_data is valid. cpr_cpu_berr 1 In CPR bus error signal to the CPU. cpr_cpu_data[31:0] 32 In Read data bus from the CPR block cpu_gpio_sel 1 Out GPIO block select. gpio_cpu_rdy 1 In GPIO ready signal to the CPU. gpio_cpu_berr 1 In GPIO bus error signal to the CPU. gpio_cpu_data[31:0] 32 In Read data bus from the GPIO block cpu_icu_sel 1 Out ICU block select. icu_cpu_rdy 1 In ICU ready signal to the CPU. icu_cpu_berr 1 In ICU bus error signal to the CPU. icu_cpu_data[31:0] 32 In Read data bus from the ICU block cpu_lss_sel 1 Out LSS block select. lss_cpu_rdy 1 In LSS ready signal to the CPU. lss_cpu_berr 1 In LSS bus error signal to the CPU. lss_cpu_data[31:0] 32 In Read data bus from the LSS block cpu_pcu_sel 1 Out PCU block select. pcu_cpu_rdy 1 In PCU ready signal to the CPU. pcu_cpu_berr 1 In PCU bus error signal to the CPU. pcu_cpu_data[31:0] 32 In Read data bus from the PCU block cpu_scb_sel 1 Out SCB block select. scb_cpu_rdy 1 In SCB ready signal to the CPU. scb_cpu_berr 1 In SCB bus error signal to the CPU. scb_cpu_data[31:0] 32 In Read data bus from the SCB block cpu_tim_sel 1 Out Timers block select. tim_cpu_rdy 1 In Timers block ready signal to the CPU. tim_cpu_berr 1 In Timers bus error signal to the CPU. tim_cpu_data[31:0] 32 In Read data bus from the Timers block cpu_rom_sel 1 Out ROM block select. rom_cpu_rdy 1 In ROM block ready signal to the CPU. rom_cpu_berr 1 In ROM bus error signal to the CPU. rom_cpu_data[31:0] 32 In Read data bus from the ROM block cpu_pss_sel 1 Out PSS block select. pss_cpu_rdy 1 In PSS block ready signal to the CPU. pss_cpu_berr 1 In PSS bus error signal to the CPU. pss_cpu_data[31:0] 32 In Read data bus from the PSS block cpu_diu_sel 1 Out DIU register block select. diu_cpu_rdy 1 In DIU register block ready signal to the CPU. diu_cpu_berr 1 In DIU bus error signal to the CPU. diu_cpu_data[31:0] 32 In Read data bus from the DIU block Interrupt signals icu_cpu_ilevel[3:0] 3 In An interrupt is asserted by driving the appropriate priority level on icu_cpu_ilevel. These signals must remain asserted until the CPU executes an interrupt acknowledge cycle. 3 Out Indicates the level of the interrupt the CPU is acknowledging when cpu_iack is high cpu_iack 1 Out Interrupt acknowledge signal. The exact timing depends on the CPU core implementation Debug signals diu_cpu_debug_valid 1 In Signal indicating the data on the diu_cpu_data bus is valid debug data. tim_cpu_debug_valid 1 In Signal indicating the data on the tim_cpu_data bus is valid debug data. scb_cpu_debug_valid 1 In Signal indicating the data on the scb_cpu_data bus is valid debug data. pcu_cpu_debug_valid 1 In Signal indicating the data on the pcu_cpu_data bus is valid debug data. lss_cpu_debug_valid 1 In Signal indicating the data on the lss_cpu_data bus is valid debug data. icu_cpu_debug_valid 1 In Signal indicating the data on the icu_cpu_data bus is valid debug data. gpio_cpu_debug_valid 1 In Signal indicating the data on the gpio_cpu_data bus is valid debug data. cpr_cpu_debug_valid 1 In Signal indicating the data on the cpr_cpu_data bus is valid debug data. debug_data_out 32 Out Output debug data to be muxed on to the GPIO & PHI pins debug_data_valid 1 Out Debug valid signal indicating the validity of the data on debug_data_out. This signal is used in all debug configurations debug_cntrl 33 Out Control signal for each PHI bound debug data line indicating whether or not the debug data should be selected by the pin mux 11.3 Realtime Requirements

The SoPEC realtime requirements have yet to be fully determined but they may be split into three categories: hard, firm and soft

11.3.1 Hard Realtime Requirements

Hard requirements are tasks that must be completed before a certain deadline or failure to do so will result in an error perceptible to the user (printing stops or functions incorrectly). There are three hard realtime tasks:

-   -   Motor control: The motors which feed the paper through the         printer at a constant speed during printing are driven directly         by the SoPEC device. Four periodic signals with different phase         relationships need to be generated to ensure the paper travels         smoothly through the printer. The generation of these signals is         handled by the GPIO hardware (see section 13.2 for more details)         but the CPU is responsible for enabling these signals (i.e. to         start or stop the motors) and coordinating the movement of the         paper with the printing operation of the printhead.     -   Buffer management: Data enters the SoPEC via the SCB at an         uneven rate and is consumed by the PEP subsystem at a different         rate. The CPU is responsible for managing the DRAM buffers to         ensure that neither overrun nor underrun occur. This buffer         management is likely to be performed under the direction of the         host.     -   Band processing: In certain cases PEP registers may need to be         updated between bands. As the timing requirements are most         likely too stringent to be met by direct CPU writes to the PCU a         more likely scenario is that a set of shadow registers will         programmed in the compressed page units before the current band         is finished, copied to band related registers by the finished         band signals and the processing of the next band will continue         immediately. An alternative solution is that the CPU will         construct a DRAM based set of commands (see section 21.8.5 for         more details) that can be executed by the PCU. The task for the         CPU here is to parse the band headers stored in DRAM and         generate a DRAM based set of commands for the next number of         bands. The location of the DRAM based set of commands must then         be written to the PCU before the current band has been processed         by the PEP subsystem. It is also conceivable (but currently         considered unlikely) that the host PC could create the DRAM         based commands. In this case the CPU will only be required to         point the PCU to the correct location in DRAM to execute         commands from.         11.3.2 Firm Requirements

Firm requirements are tasks that should be completed by a certain time or failure to do so will result in a degradation of performance but not an error. The majority of the CPU tasks for SoPEC fall into this category including all interactions with the QA chips, program authentication, page feeding, configuring PEP registers for a page or job, determining the firing pulse profile, communication of printer status to the host over the USB and the monitoring of ink usage. The authentication of downloaded programs and messages will be the most compute intensive operation the CPU will be required to perform. Initial investigations indicate that the LEON processor, running at 160 MHz, will easily perform three authentications in under a second.

TABLE 15 Expected firm requirements Requirement Duration Power-on to start of printing first ~8 secs ?? page [USB and slave SoPEC enumeration, 3 or more RSA signature verifications, code and compressed page data download and chip initialisation] Wake-up from sleep mode to start ~2 secs printing [3 or more SHA-1/RSA operations, code and compressed page data download and chip re- initialisation Authenticate ink usage in the printer ~0.5 secs Determining firing pulse profile ~0.1 secs Page feeding, gap between pages OEM dependent Communication of printer status ~10 ms to host PC Configuring PEP registers ?? 11.3.3 Soft Requirements

Soft requirements are tasks that need to be done but there are only light time constraints on when they need to be done. These tasks are performed by the CPU when there are no pending higher priority tasks. As the SoPEC CPU is expected to be lightly loaded these tasks will mostly be executed soon after they are scheduled.

11.4 Bus Protocols

As can be seen from FIG. 15 above there are different buses in the CPU block and different protocols are used for each bus. There are three buses in operation:

11.4.1 AHB Bus

The LEON CPU core uses an AMBA2.0 AHB bus to communicate with memory and peripherals (usually via an APB bridge). See the AMBA specification [38], section 5 of the LEON users manual [37] and section 11.6.6.1 of this document for more details.

11.4.2 CPU to DIU Bus

This bus conforms to the DIU bus protocol described in Section 20.14.8. Note that the address bus used for DIU reads (i.e. cpu_adr(21:2)) is also that used for CPU subsystem with bus accesses while the write address bus (cpu_diu_wadr) and the read and write data buses (dram_cpu_data and cpu_diu_wdata) are private buses between the CPU and the DIU. The effective bus width differs between a read (256 bits) and a write (128 bits). As certain CPU instructions may require byte write access this will need to be supported by both the DRAM write buffer (in the AHB bridge) and the DIU. See section 11.6.6.1 for more details.

11.4.3 CPU Subsystem Bus

For access to the on-chip peripherals a simple bus protocol is used. The MMU must first determine which particular block is being addressed (and that the access is a valid one) so that the appropriate block select signal can be generated. During a write access CPU write data is driven out with the address and block select signals in the first cycle of an access. The addressed slave peripheral responds by asserting its ready signal indicating that it has registered the write data and the access can complete. The write data bus is common to all peripherals and is also used for CPU writes to the embedded DRAM. A read access is initiated by driving the address and select signals during the first cycle of an access. The addressed slave responds by placing the read data on its bus and asserting its ready signal to indicate to the CPU that the read data is valid. Each block has a separate point-to-point data bus for read accesses to avoid the need for a tri-stateable bus. All peripheral accesses are 32-bit (Programming note: char or short C types should not be used to access peripheral registers). The use of the ready signal allows the accesses to be of variable length. In most cases accesses will complete in two cycles but three or four (or more) cycles accesses are likely for PEP blocks or IP blocks with a different native bus interface. All PEP blocks are accessed via the PCU which acts as a bridge. The PCU bus uses a similar protocol to the CPU subsystem bus but with the PCU as the bus master.

The duration of accesses to the PEP blocks is influenced by whether or not the PCU is executing commands from DRAM. As these commands are essentially register writes the CPU access will need to wait until the PCU bus becomes available when a register access has been completed. This could lead to the CPU being stalled for up to 4 cycles if it attempts to access PEP blocks while the PCU is executing a command. The size and probability of this penalty is sufficiently small to have any significant impact on performance.

In order to support user mode (i.e. OEM code) access to certain peripherals the CPU subsystem bus propagates the CPU function code signals (cpu_acode[1:0]). These signals indicate the type of address space (i.e. User/Supervisor and Program/Data) being accessed by the CPU for each access. Each peripheral must determine whether or not the CPU is in the correct mode to be granted access to its registers and in some cases (e.g. Timers and GPIO blocks) different access permissions can apply to different registers within the block. If the CPU is not in the correct mode then the violation is flagged by asserting the block's bus error signal (block_cpu_berr) with the same timing as its ready signal (block_cpu_rdy) which remains deasserted. When this occurs invalid read accesses should return 0 and write accesses should have no effect.

FIG. 16 shows two examples of the peripheral bus protocol in action. A write to the LSS block from code running in supervisor mode is successfully completed. This is immediately followed by a read from a PEP block via the PCU from code running in user mode. As this type of access is not permitted the access is terminated with a bus error. The bus error exception processing then starts directly after this—no further accesses to the peripheral should be required as the exception handler should be located in the DRAM.

Each peripheral acts as a slave on the CPU subsystem bus and its behavior is described by the state machine in section 11.4.3.1

11.4.3.1 CPU Subsystem Bus Slave State Machine

CPU subsystem bus slave operation is described by the state machine in FIG. 17. This state machine will be implemented in each CPU subsystem bus slave. The only new signals mentioned here are the valid_access and reg_available signals. The valid_access is determined by comparing the cpu_acode value with the block or register (in the case of a block that allow user access on a per register basis such as the GPIO block) access permissions and asserting valid_access if the permissions agree with the CPU mode. The reg_available signal is only required in the PCU or in blocks that are not capable of two-cycle access (e.g. blocks containing imported IP with different bus protocols). In these blocks the reg_available signal is an internal signal used to insert wait states (by delaying the assertion of block_cpu_rdy) until the CPU bus slave interface can gain access to the register.

When reading from a register that is less than 32 bits wide the CPU subsystems bus slave should return zeroes on the unused upper bits of the block_cpu_data bus.

To support debug mode the contents of the register selected for debug observation, debug_reg, are always output on the block_cpu_data bus whenever a read access is not taking place. See section

11.8 for More Details of Debug Operation.

11.5 LEON CPU

The LEON processor is an open-source implementation of the IEEE-1754 standard (SPARC V8) instruction set. LEON is available from and actively supported by Gaisler Research (www.gaisler.com).

The following features of the LEON-2 processor will be utilised on SoPEC:

-   -   IEEE-1754 (SPARC V8) compatible integer unit with 5-stage         pipeline     -   Separate instruction and data cache (Harvard architecture). 1         kbyte direct mapped caches will be used for both.     -   Full implementation of AMBA-2.0 AHB on-chip bus

The standard release of LEON incorporates a number of peripherals and support blocks which will not be included on SoPEC. The LEON core as used on SoPEC will consist of: 1) the LEON integer unit, 2) the instruction and data caches (currently 1 kB each), 3) the cache control logic, 4) the AHB interface and 5) possibly the AHB controller (although this functionality may be implemented in the LEON AHB bridge).

The version of the LEON database that the SoPEC LEON components will be sourced from is LEON2-1.0.7 although later versions may be used if they offer worthwhile functionality or bug fixes that affect the SoPEC design.

The LEON core will be clocked using the system clock, pclk, and reset using the prst_n_section[1] signal. The ICU will assert all the hardware interrupts using the protocol described in section 11.9. The LEON hardware multipliers and floating-point unit are not required. SoPEC will use the recommended 8 register window configuration.

Further details of the SPARC V8 instruction set and the LEON processor can be found in [36] and [37] respectively.

11.5.1 LEON Registers

Only two of the registers described in the LEON manual are implemented on SoPEC—the LEON configuration register and the Cache Control Register (CCR). The addresses of these registers are shown in Table 16. The configuration register bit fields are described below and the CCR is described in section 11.7.1.1.

11.5.1.1 LEON Configuration Register

The LEON configuration register allows runtime software to determine the settings of LEONs various configuration options. This is a read-only register whose value for the SoPEC ASIC will be 0x1071_(—)8C00. Further descriptions of many of the bitfileds can be found in the LEON manual. The values used for SoPEC are highlighted in bold for clarity.

TABLE 16 LEON Configuration Register Field Name bit(s) Description WriteProtection 1:0 Write protection type. 00 - none 01 - standard PCICore 3:2 PCI core type 00 - none 01 - InSilicon 10 - ESA 11 - Other FPUType 5:4 FPU type. 00 - none 01 - Meiko MemStatus 6 0 - No memory status and failing address register present 1 - Memory status and failing address register present Watchdog 7 0 - Watchdog timer not present (Note this refers to the LEON watchdog timer in the LEON timer block). 1 - Watchdog timer present UMUL/SMUL 8 0 - UMUL/SMUL instructions are not implemented 1 - UMUL/SMUL instructions are implemented UDIV/SDIV 9 0 - UMUL/SMUL instructions are not implemented 1 - UMUL/SMUL instructions are implemented DLSZ 11:10 Data cache line size in 32-bit words: 00 - 1 word 01 - 2 words 10 - 4 words 11 - 8 words DCSZ 14:12 Data cache size in kBbytes = 2^(DCSZ). SoPEC DCSZ = 0. ILSZ 16:15 Instruction cache line size in 32-bit words: 00 - 1 word 01 - 2 words 10 - 4 words 11 - 8 words ICSZ 19:17 Instruction cache size in kBbytes = 2^(ICSZ). SoPEC ICSZ = 0. RegWin 24:20 The implemented number of SPARC register windows − 1. SoPEC value = 7. UMAC/SMAC 25 0 - UMAC/SMAC instructions are not implemented 1 - UMAC/SMAC instructions are implemented Watchpoints 28:26 The implemented number of hardware watchpoints. SoPEC value = 4. SDRAM 29 0 - SDRAM controller not present 1 - SDRAM controller present DSU 30 0 - Debug Support Unit not present 1 - Debug Support Unit present Reserved 31 Reserved. SoPEC value = 0. 11.6 Memory Management Unit (MMU)

Memory Management Units are typically used to protect certain regions of memory from invalid accesses, to perform address translation for a virtual memory system and to maintain memory page status (swapped-in, swapped-out or unmapped)

The SoPEC MMU is a much simpler affair whose function is to ensure that all regions of the SoPEC memory map are adequately protected. The MMU does not support virtual memory and physical addresses are used at all times. The SoPEC MMU supports a full 32-bit address space. The SoPEC memory map is depicted in FIG. 18 below.

The MMU selects the relevant bus protocol and generates the appropriate control signals depending on the area of memory being accessed. The MMU is responsible for performing the address decode and generation of the appropriate block select signal as well as the selection of the correct block read bus during a read access. The MMU will need to support all of the bus transactions the CPU can produce including interrupt acknowledge cycles, aborted transactions etc. When an MMU error occurs (such as an attempt to access a supervisor mode only region when in user mode) a bus error is generated. While the LEON can recognise different types of bus error (e.g. data store error, instruction access error) it handles them in the same manner as it handles all traps i.e it will transfer control to a trap handler. No extra state information is be stored because of the nature of the trap. The location of the trap handler is contained in the TBR (Trap Base Register). This is the same mechanism as is used to handle interrupts.

11.6.1 CPU-Bus Peripherals Address Map

The address mapping for the peripherals attached to the CPU-bus is shown in Table 17 below. The MMU performs the decode of the high order bits to generate the relevant cpu_block_select signal. Apart from the PCU, which decodes the address space for the PEP blocks, each block only needs to decode as many bits of cpu_adr[1 1:2] as required to address all the registers within the block.

TABLE 17 CPU-bus peripherals address map Block_base Address ROM_base 0x0000_0000 MMU_base 0x0001_0000 TIM_base 0x0001_1000 LSS_base 0x0001_2000 GPIO_base 0x0001_3000 SCB_base 0x0001_4000 ICU_base 0x0001_5000 CPR_base 0x0001_6000 DIU_base 0x0001_7000 PSS_base 0x0001_8000 Reserved 0x0001_9000 to 0x0001_FFFF PCU_base 0x0002_0000 11.6.2 DRAM Region Mapping

The embedded DRAM is broken into 8 regions, with each region defined by a lower and upper bound address and with its own access permissions.

The association of an area in the DRAM address space with a MMU region is completely under software control. Table 18 below gives one possible region mapping. Regions should be defined according to their access requirements and position in memory. Regions that share the same access requirements and that are contiguous in memory may be combined into a single region. The example below is purely for indicative purposes—real mappings are likely to differ significantly from this. Note that the RegionBottom and RegionTop fields in this example include the DRAM base address offset (0x4000_(—)0000) which is not required when programming the RegionNTop and RegionNBottom registers. For more details, see 11.6.5.1 and 11.6.5.2.

TABLE 18 Example region mapping Region RegionBottom RegionTop Description 0 0x4000_0000 0x4000_0FFF Silverbrook OS (supervisor) data 1 0x4000_1000 0x4000_BFFF Silverbrook OS (supervisor) code 2 0x4000_C000 0x4000_C3FF Silverbrook (supervisor/user) data 3 0x4000_C400 0x4000_CFFF Silverbrook (supervisor/user) code 4 0x4026_D000 0x4026_D3FF OEM (user) data 5 0x4026_D400 0x4026_DFFF OEM (user) code 6 0x4027_E000 0x4027_FFFF Shared Silverbrook/ OEM space 7 0x4000_D000 0x4026_CFFF Compressed page store (supervisor data) 11.6.3 Non-DRAM Regions

As shown in FIG. 18 the DRAM occupies only 2.5 MBytes of the total 4 GB SoPEC address space. The non-DRAM regions of SoPEC are handled by the MMU as follows:

ROM (1x0000_(—)000 to 1x0000_FFFF): The ROM block will control the access types allowed. The cpu_acode[1:0] signals will indicate the CPU mode and access type and the ROM block will assert rom_cpu_berr if an attempted access is forbidden. The protocol is described in more detail in section 11.4.3. The ROM block access permissions are hard wired to allow all read accesses except to the FuseChipID registers which may only be read in supervisor mode.

MMU Internal Registers (0x0001_(—)0000 to 1x0001_(—)0FFF): The MMU is responsible for controlling the accesses to its own internal registers and will only allow data reads and writes (no instruction fetches) from supervisor data space. All other accesses will result in the mmu_cpu_berr signal being asserted in accordance with the CPU native bus protocol.

CPU Subsystem Peripheral Registers (1x0001_(—)1000 to 1x0001_FFFF): Each peripheral block will control the access types allowed. Every peripheral will allow supervisor data accesses (both read and write) and some blocks (e.g. Timers and GPIO) will also allow user data space accesses as outlined in the relevant chapters of this specification. Neither supervisor nor user instruction fetch accesses are allowed to any block as it is not possible to execute code from peripheral registers. The bus protocol is described in section 11.4.3.

PCU Mapped Registers (0x0002_(—)0000 to 0x0002_BFFF): All of the PEP blocks registers which are accessed by the CPU via the PCU will inherit the access permissions of the PCU. These access permissions are hard wired to allow supervisor data accesses only and the protocol used is the same as for the CPU peripherals.

Unused address space (0x0002_C000 to 0x3FFF_FFFF and 0x4028_(—)0000 to 1xFFFF_FFFF): All accesses to the unused portion of the address space will result in the mmu_cpu_berr signal being asserted in accordance with the CPU native bus protocol. These accesses will not propagate outside of the MMU i.e. no external access will be initiated.

11.6.4 Reset Exception Vector and Reference Zero Traps

When a reset occurs the LEON processor starts executing code from address 0x0000_(—)0000. A common software bug is zero-referencing or null pointer de-referencing (where the program attempts to access the contents of address 0x0000_(—)0000). To assist software debug the MMU will assert a bus error every time the locations 0x0000_(—)0000 to 0x0000_(—)000F (i.e. the first 4 words of the reset trap) are accessed after the reset trap handler has legitimately been retrieved immediately after reset.

11.6.5 MMU Configuration Registers

The MMU configuration registers include the RDU configuration registers and two LEON registers. Note that all the MMU configuration registers may only be accessed when the CPU is running in supervisor mode.

TABLE 19 MMU Configuration Registers Address offset from MMU_base Register #bits Reset Description 0x00 Region0Bottom[21:5] 17 0x0_0000 This register contains the physical address that marks the bottom of region 0 0x04 Region0Top[21:5] 17 0xF_FFFF This register contains the physical address that marks the top of region 0. Region 0 covers the entire address space after reset whereas all other regions are zero-sized initially. 0x08 Region1Bottom[21:5] 17 0xF_FFFF This register contains the physical address that marks the bottom of region 1 0x0C Region1Top[21:5] 17 0x0_0000 This register contains the physical address that marks the top of region 1 0x10 Region2Bottom[21:5] 17 0xF_FFFF This register contains the physical address that marks the bottom of region 2 0x14 Region3Top[21:5] 17 0x0_0000 This register contains the physical address that marks the top of region 2 0x18 Region3Bottom[21:5] 17 0xF_FFFF This register contains the physical address that marks the bottom of region 3 0x1C Region3Top[21:5] 17 0x0_0000 This register contains the physical address that marks the top of region 3 0x20 Region4Bottom[21:5] 17 0xF_FFFF This register contains the physical address that marks the bottom of region 4 0x24 Region4Top[21:5] 17 0x0_0000 This register contains the physical address that marks the top of region 4 0x28 Region5Bottom[21:5] 17 0xF_FFFF This register contains the physical address that marks the bottom of region 5 0x2C Region5Top[21:5] 17 0x0_0000 This register contains the physical address that marks the top of region 5 0x30 Region6Bottom[21:5] 17 0xF_FFFF This register contains the physical address that marks the bottom of region 6 0x34 Region6Top[21:5] 17 0x0_0000 This register contains the physical address that marks the top of region 6 0x38 Region7Bottom[21:5] 17 0xF_FFFF This register contains the physical address that marks the bottom of region 7 0x3C Region7Top[21:5] 17 0x0_0000 This register contains the physical address that marks the top of region 7 0x40 Region0Control 6 0x07 Control register for region 0 0x44 Region1Control 6 0x07 Control register for region 1 0x48 Region2Control 6 0x07 Control register for region 2 0x4C Region3Control 6 0x07 Control register for region 3 0x50 Region4Control 6 0x07 Control register for region 4 0x54 Region5Control 6 0x07 Control register for region 5 0x58 Region6Control 6 0x07 Control register for region 6 0x5C Region7Control 6 0x07 Control register for region 7 0x60 RegionLock 8 0x00 Writing a 1 to a bit in the RegionLock register locks the value of the corresponding RegionTop, RegionBottom and RegionControl regis- ters. The lock can only be cleared by a reset and any attempt to write to a locked register will result in a bus error. 0x64 BusTimeout 8 0xFF This register should be set to the number of pclk cycles to wait after an access has started before aborting the access with a bus error. Writing 0 to this register disables the bus time- out feature. 0x68 ExceptionSource 6 0x00 This register identifies the source of the last exception. See Section 11.6.5.3 for details. 0x6C DebugSelect 7 0x00 Contains address of the register selected for debug observation. It is expected that a number of pseudo-registers will be made available for debug observation and these will be outlined during the implementation phase. 0x80 to RDU Registers See Table for details. 0x108 0x140 LEON Configuration 32 0x1071_8C00 The LEON configuration register is used by Register software to determine the configuration of this LEON implementation. See section 11.5.1.1 for details. This register is ReadOnly. 0x144 LEON Cache 32 0x0000_0000 The LEON Cache Control Register is used to Control Register control the operation of the caches. See section 11.6 for details. 11.6.5.1 RegionTop and RegionBottom Registers

The 20 Mbit of embedded DRAM on SoPEC is arranged as 81920 words of 256 bits each. All region boundaries need to align with a 256-bit word. Thus only 17 bits are required for the RegionNTop and RegionNBottom registers. Note that the bottom 5 bits of the RegionNTop and RegionNBottom registers cannot be written to and read as ‘0’ i.e. the RegionNTop and RegionNBottom registers represent byte-aligned DRAM addresses

Both the RegionNTop and RegionNBottom registers are inclusive i.e. the addresses in the registers are included in the region. Thus the size of a region is (RegionNTop—RegionNBottom)+1 DRAM words.

If DRAM regions overlap (there is no reason for this to be the case but there is nothing to prohibit it either) then only accesses allowed by all overlapping regions are permitted. That is if a DRAM address appears in both Region1 and Region3 (for example) the cpu_acode of an access is checked against the access permissions of both regions. If both regions permit the access then it will proceed but if either or both regions do not permit the access then it will not be allowed. The MMU does not support negatively sized regions i.e. the value of the RegionNTop register should always be greater than or equal to the value of the RegionNBottom register. If RegionNTop is lower in the address map than RegionNTop then the region is considered to be zero-sized and is ignored.

When both the RegionNTop and RegionNBottom registers for a region contain the same value the region is then simply one 256-bit word in length and this corresponds to the smallest possible active region.

11.6.5.2 Region Control Registers

Each memory region has a control register associated with it. The RegionNControl register is used to set the access conditions for the memory region bounded by the RegionNTop and RegionNBottom registers. Table 20 describes the function of each bit field in the RegionNControl registers. All bits in a RegionNControl register are both readable and writable by design. However, like all registers in the MMU, the RegionNControl registers can only be accessed by code running in supervisor mode.

TABLE 20 Region Control Register Field Name bit(s) Description SupervisorAccess 2:0 Denotes the type of access allowed when the CPU is running in Supervisor mode. For each access type a 1 indicates the access is per- mitted and a 0 indicates the access is not permitted. bit0 - Data read access permission bit1 - Data write access permission bit2 - Instruction fetch access permission UserAccess 5:3 Denotes the type of access allowed when the CPU is running in User mode. For each access type a 1 indicates the access is permitted and a 0 indicates the access is not permitted. bit3 - Data read access permission bit4 - Data write access permission bit5 - Instruction fetch access permission 11.6.5.3 ExceptionSource Register

The SPARC V8 architecture allows for a number of types of memory access error to be trapped. These trap types and trap handling in general are described in chapter 7 of the SPARC architecture manual [36]. However on the LEON processor only data_store_error and data_access_exception trap types will result from an external (to LEON) bus error. According to the SPARC architecture manual the processor will automatically move to the next register window (i.e. it decrements the current window pointer) and copies the program counters (PC and nPC) to two local registers in the new window. The supervisor bit in the PSR is also set and the PSR can be saved to another local register by the trap handler (this does not happen automatically in hardware). The ExceptionSource register aids the trap handler by identifying the source of an exception. Each bit in the ExceptionSource register is set when the relevant trap condition and should be cleared by the trap handler by writing a ‘1’ to that bit position.

TABLE 21 ExceptionSource Register Field Name bit(s) Description DramAccessExcptn 0 The permissions of an access did not match those of the DRAM region it was attempting to access. This bit will also be set if an attempt is made to access an undefined DRAM region (i.e. a location that is not within the bounds of any RegionTop/RegionBottom pair) PeriAccessExcptn 1 An access violation occurred when accessing a CPU subsystem block. This occurs when the access per- missions disagree with those set by the block. UnusedAreaExcptn 2 An attempt was made to access an unused part of the memory map LockedWriteExcptn 3 An attempt was made to write to a regions registers (RegionTop/ Bottom/Control) after they had been locked. ResetHandlerExcptn 4 An attempt was made to access a ROM location between 0x0000_0000 and 0x0000_000F after the reset handler was executed. The most likely cause of such an access is the use of an uninitialised pointer or structure. TimeoutExcptn 5 A bus timeout condition occurred. 11.6.6 MMU Sub-Block Partition

As can be seen from FIG. 19 and FIG. 20 the MMU consists of three principal sub-blocks. For clarity the connections between these sub-blocks and other SoPEC blocks and between each of the sub-blocks are shown in two separate diagrams.

11.6.6.1 LEON AHB Bridge

The LEON AHB bridge consists of an AHB bridge to DIU and an AHB to CPU subsystem bus bridge. The AHB bridge will convert between the AHB and the DIU and CPU subsystem bus protocols but the address decoding and enabling of an access happens elsewhere in the MMU. The AHB bridge will always be a slave on the AHB. Note that the AMBA signals from the LEON core are contained within the ahbso and ahbsi records. The LEON records are described in more detail in section 11.7. Glue logic may be required to assist with enabling memory accesses, endianness coherency, interrupts and other miscellaneous signalling.

TABLE 22 LEON AHB bridge I/Os Port name Pins I/O Description Global SoPEC signals prst_n 1 In Global reset. Synchronous to pclk, active low. pclk 1 In Global clock LEON core to LEON AHB signals (ahbsi and ahbso records) ahbsi.haddr[31:0] 32 In AHB address bus ahbsi.hwdata[31:0] 32 In AHB write data bus ahbso.hrdata[31:0] 32 Out AHB read data bus ahbsi.hsel 1 In AHB slave select signal ahbsi.hwrite 1 In AHB write signal: 1 - Write access 0 - Read access ahbsi.htrans 2 In Indicates the type of the current transfer: 00 - IDLE 01 - BUSY 10 - NONSEQ 11 - SEQ ahbsi.hsize 3 In Indicates the size of the current transfer: 000 - Byte transfer 001 - Halfword transfer 010 - Word transfer 011 - 64-bit transfer (unsupported?) 1xx - Unsupported larger wordsizes ahbsi.hburst 3 In Indicates if the current transfer forms part of a burst and the type of burst: 000 - SINGLE 001 - INCR 010 - WRAP4 011 - INCR4 100 - WRAP8 101 - INCR8 110 - WRAP16 111 - INCR16 ahbsi.hprot 4 In Protection control signals pertaining to the current access: hprot[0] - Opcode(0)/Data(1) access hprot[1] - User(0)/Supervisor access hprot[2] - Non-bufferable(0)/Bufferable(1) access (unsupported) hprot[3] - Non-cacheable(0)/Cacheable access ahbsi.hmaster 4 In Indicates the identity of the current bus master. This will always be the LEON core. ahbsi.hmastlock 1 In Indicates that the current master is performing a locked sequence of transfers. ahbso.hready 1 Out Active high ready signal indicating the access has completed ahbso.hresp 2 Out Indicates the status of the transfer: 00 - OKAY 01 - ERROR 10 - RETRY 11 - SPLIT ahbso.hsplit[15:0] 16 Out This 16-bit split bus is used by a slave to indicate to the arbiter which bus masters should be allowed attempt a split transaction. This feature will be unsupported on the AHB bridge Toplevel/Common LEON AHB bridge signals cpu_dataout[31:0] 32 Out Data out bus to both DRAM and peripheral devices. cpu_rwn 1 Out Read/NotWrite signal. 1 = Current access is a read access, 0 = Current access is a write access icu_cpu_ilevel[3:0] 4 In An interrupt is asserted by driving the appropriate priority level on icu_cpu_ilevel. These signals must remain asserted until the CPU executes an interrupt acknowledge cycle. cpu_icu_ilevel[3:0] 4 In Indicates the level of the interrupt the CPU is acknowledging when cpu_iack is high cpu_iack 1 Out Interrupt acknowledge signal. The exact timing depends on the CPU core implementation cpu_start_access 1 Out Start Access signal indicating the start of a data transfer and that the cpu_adr, cpu_dataout, cpu_rwn and cpu_acode signals are all valid. This signal is only asserted during the first cycle of an access. cpu_ben[1:0] 2 Out Byte enable signals. dram_cpu_data[255:0] 256 In Read data from the DRAM. diu_cpu_rreq 1 Out Read request to the DIU. diu_cpu_rack 1 In Acknowledge from DIU that read request has been accepted. diu_cpu_rvalid 1 In Signal from DIU indicating that valid read data is on the dram_cpu_data bus cpu_diu_wdatavalid 1 Out Signal from the CPU to the DIU indicating that the data currently on the cpu_diu_wdata bus is valid and should be committed to the DIU posted write buffer diu_cpu_write_rdy 1 In Signal from the DIU indicating that the posted write buffer is empty cpu_diu_wdadr[21:4] 18 Out Write address bus to the DIU cpu_diu_wdata[127:0] 128 Out Write data bus to the DIU cpu_diu_wmask[15:0] 16 Out Write mask for the cpu_diu_wdata bus. Each bit corresponds to a byte of the 128-bit cpu_diu_wdata bus. LEON AHB bridge to MMU Control Block signals cpu_mmu_adr 32 Out CPU Address Bus. mmu_cpu_data 32 In Data bus from the MMU mmu_cpu_rdy 1 In Ready signal from the MMU cpu_mmu_acode 2 Out Access code signals to the MMU mmu_cpu_berr 1 In Bus error signal from the MMU dram_access_en 1 In DRAM access enable signal. A DRAM access cannot be initiated unless it has been enabled by the MMU control unit. Description:

The LEON AHB bridge must ensure that all CPU bus transactions are functionally correct and that the timing requirements are met. The AHB bridge also implements a 128-bit DRAM write buffer to improve the efficiency of DRAM writes, particularly for multiple successive writes to DRAM. The AHB bridge is also responsible for ensuring endianness coherency i.e. guaranteeing that the correct data appears in the correct position on the data buses (hrdata, cpu_dataout and cpu_mmu_wdata) for every type of access. This is a requirement because the LEON uses big-endian addressing while the rest of SoPEC is little-endian.

The LEON AHB bridge will assert request signals to the DIU if the MMU control block deems the access to be a legal access. The validity (i.e. is the CPU running in the correct mode for the address space being accessed) of an access is determined by the contents of the relevant RegionNControl register. As the SPARC standard requires that all accesses are aligned to their word size (i.e. byte, half-word, word or double-word) and so it is not possible for an access to traverse a 256-bit boundary (as required by the DIU). Invalid DRAM accesses are not propagated to the DIU and will result in an error response (ahbso.hresp=‘01’) on the AHB. The DIU bus protocol is described in more detail in section 20.9. The DIU will return a 256-bit dataword on dram_cpu_data[255:0] for every read access.

The CPU subsystem bus protocol is described in section 11.4.3. While the LEON AHB bridge performs the protocol translation between AHB and the CPU subsystem bus the select signals for each block are generated by address decoding in the CPU subsystem bus interface. The CPU subsystem bus interface also selects the correct read data bus, ready and error signals for the block being addressed and passes these to the LEON AHB bridge which puts them on the AHB bus. It is expected that some signals (especially those external to the CPU block) will need to be registered here to meet the timing requirements. Careful thought will be required to ensure that overall CPU access times are not excessively degraded by the use of too many register stages.

11.6.6.1.1 DRAM Write Buffer

The DRAM write buffer improves the efficiency of DRAM writes by aggregating a number of CPU write accesses into a single DIU write access. This is achieved by checking to see if a CPU write is to an address already in the write buffer and if so the write is immediately acknowledged (i.e. the ahbsi.hready signal is asserted without any wait states) and the DRAM write buffer updated accordingly. When the CPU write is to a DRAM address other than that in the write buffer then the current contents of the write buffer are sent to the DIU (where they are placed in the posted write buffer) and the DRAM write buffer is updated with the address and data of the CPU write. The DRAM write buffer consists of a 128-bit data buffer, an 18-bit write address tag and a 16-bit write mask. Each bit of the write mask indicates the validity of the corresponding byte of the write buffer as shown in FIG. 21 below.

The operation of the DRAM write buffer is summarised by the following set of rules:

-   1) The DRAM write buffer only contains DRAM write data i.e.     peripheral writes go directly to the addressed peripheral. -   2) CPU writes to locations within the DRAM write buffer or to an     empty write buffer (i.e. the write mask bits are all 0) complete     with zero wait states regardless of the size of the write     (byte/half-word/word/ double-word). -   3) The contents of the DRAM write buffer are flushed to DRAM     whenever a CPU write to a location outside the write buffer occurs,     whenever a CPU read from a location within the write buffer occurs     or whenever a write to a peripheral register occurs. -   4) A flush resulting from a peripheral write will not cause any     extra wait states to be inserted in the peripheral write access. -   5) Flushes resulting from a DRAM accesses will cause wait states to     be inserted until the DIU posted write buffer is empty. If the DIU     posted write buffer is empty at the time the flush is required then     no wait states will be inserted for a flush resulting from a CPU     write or one wait state will be inserted for a flush resulting from     a CPU read (this is to ensure that the DIU sees the write request     ahead of the read request). Note that in this case further wait     states will also be inserted as a result of the delay in servicing     the read request by the DIU.     11.6.6.1.2 DIU Interface Waveforms

FIG. 22 below depicts the operation of the AHB bridge over a sample sequence of DRAM transactions consisting of a read into the DCache, a double-word store to an address other than that currently in the DRAM write buffer followed by an ICache line refill. To avoid clutter a number of AHB control signals that are inputs to the MMU have been grouped together as ahbsi.CONTROL and only the ahbso.HREADY is shown of the output AHB control signals.

The first transaction is a single word load (‘LD’). The MMU (specifically the MMU control block) uses the first cycle of every access (i.e. the address phase of an AHB transaction) to determine whether or not the access is a legal access. The read request to the DIU is then asserted in the following cycle (assuming the access is a valid one) and is acknowledged by the DIU a cycle later. Note that the time from cpu_diu_rreq being asserted and diu_cpu_rack being asserted is variable as it depends on the DIU configuration and access patterns of DIU requestors. The AHB bridge will insert wait states until it sees the diu_cpu_rvalid signal is high, indicating the data (‘LD1’) on the dram_cpu_data bus is valid. The AHB bridge terminates the read access in the same cycle by asserting the ahbso.HREADY signal (together with an ‘OKAY’ HRESP code). The AHB bridge also selects the appropriate 32 bits (‘RD1’) from the 256-bit DRAM line data (‘LD1’) returned by the DIU corresponding to the word address given by A1.

The second transaction is an AHB two-beat incrementing burst issued by the LEON acache block in response to the execution of a double-word store instruction. As LEON is a big endian processor the address issued (‘A2’) during the address phase of the first beat of this transaction is the address of the most significant word of the double-word while the address for the second beat (‘A3’) is that of the least significant word i.e. A3=A2+4. The presence of the DRAM write buffer allows these writes to complete without the insertion of any wait states. This is true even when, as shown here, the DRAM write buffer needs to be flushed into the DIU posted write buffer, provided the DIU posted write buffer is empty. If the DIU posted write buffer is not empty (as would be signified by diu_cpu_write_rdy being low) then wait states would be inserted until it became empty. The cpu_diu_wdata buffer builds up the data to be written to the DIU over a number of transactions (‘BD1’ and ‘BD2’ here) while the cpu_dui_wmask records every byte that has been written to since the last flush—in this case the lowest word and then the second lowest word are written to as a result of the double-word store operation.

The final transaction shown here is a DRAM read caused by an ICache miss. Note that the pipelined nature of the AHB bus allows the address phase of this transaction to overlap with the final data phase of the previous transaction. All ICache misses appear as single word loads (‘LD’) on the AHB bus. In this case we can see that the DIU is slower to respond to this read request than to the first read request because it is processing the write access caused by the DRAM write buffer flush. The ICache refill will complete just after the window shown in FIG. 22.

11.6.6.2 CPU Subsystem Bus Interface

The CPU Subsystem Interface block handles all valid accesses to the peripheral blocks that comprise the CPU Subsystem.

TABLE 23 CPU Subsystem Bus Interface I/Os Port name Pins I/O Description Global SoPEC signals prst_n 1 In Global reset. Synchronous to pclk, active low. pclk 1 In Global clock Toplevel/Common CPU Subsystem Bus Interface signals cpu_cpr_sel 1 Out CPR block select. cpu_gpio_sel 1 Out GPIO block select. cpu_icu_sel 1 Out ICU block select. cpu_lss_sel 1 Out LSS block select. cpu_pcu_sel 1 Out PCU block select. cpu_scb_sel 1 Out SCB block select. cpu_tim_sel 1 Out Timers block select. cpu_rom_sel 1 Out ROM block select. cpu_pss_sel 1 Out PSS block select. cpu_diu_sel 1 Out DIU block select. cpr_cpu_data[31:0] 32 In Read data bus from the CPR block gpio_cpu_data[31:0] 32 In Read data bus from the GPIO block icu_cpu_data[31:0] 32 In Read data bus from the ICU block lss_cpu_data[31:0] 32 In Read data bus from the LSS block pcu_cpu_data[31:0] 32 In Read data bus from the PCU block scb_cpu_data[31:0] 32 In Read data bus from the SCB block tim_cpu_data[31:0] 32 In Read data bus from the Timers block rom_cpu_data[31:0] 32 In Read data bus from the ROM block pss_cpu_data[31:0] 32 In Read data bus from the PSS block diu_cpu_data[31:0] 32 In Read data bus from the DIU block cpr_cpu_rdy 1 In Ready signal to the CPU. When cpr_cpu_rdy is high it indicates the last cycle of the access. For a write cycle this means cpu_dataout has been registered by the CPR block and for a read cycle this means the data on cpr_cpu_data is valid. gpio_cpu_rdy 1 In GPIO ready signal to the CPU. icu_cpu_rdy 1 In ICU ready signal to the CPU. lss_cpu_rdy 1 In LSS ready signal to the CPU. pcu_cpu_rdy 1 In PCU ready signal to the CPU. scb_cpu_rdy 1 In SCB ready signal to the CPU. tim_cpu_rdy 1 In Timers block ready signal to the CPU. rom_cpu_rdy 1 In ROM block ready signal to the CPU. pss_cpu_rdy 1 In PSS block ready signal to the CPU. diu_cpu_rdy 1 In DIU register block ready signal to the CPU. cpr_cpu_berr 1 In Bus Error signal from the CPR block gpio_cpu_berr 1 In Bus Error signal from the GPIO block icu_cpu_berr 1 In Bus Error signal from the ICU block lss_cpu_berr 1 In Bus Error signal from the LSS block pcu_cpu_berr 1 In Bus Error signal from the PCU block scb_cpu_berr 1 In Bus Error signal from the SCB block tim_cpu_berr 1 In Bus Error signal from the Timers block rom_cpu_berr 1 In Bus Error signal from the ROM block pss_cpu_berr 1 In Bus Error signal from the PSS block diu_cpu_berr 1 In Bus Error signal from the DIU block CPU Subsystem Bus Interface to MMU Control Block signals cpu_adr[19:12] 8 In Toplevel CPU Address bus. Only bits 19-12 are required to decode the peripherals address space peri_access_en 1 In Enable Access signal. A peripheral access cannot be initiated unless it has been enabled by the MMU Control Unit peri_mmu_data[31:0] 32 Out Data bus from the selected peripheral peri_mmu_rdy 1 Out Data Ready signal. Indicates the data on the peri_mmu_data bus is valid for a read cycle or that the data was successfully written to the peripheral for a write cycle. peri_mmu_berr 1 Out Bus Error signal. Indicates a bus error has occurred in accessing the selected peripheral CPU Subsystem Bus Interface to LEON AHB bridge signals cpu_start_access 1 In Start Access signal from the LEON AHB bridge indicating the start of a data transfer and that the cpu_adr, cpu_dataout, cpu_rwn and cpu_acode signals are all valid. This signal is only asserted during the first cycle of an access. Description:

The CPU Subsystem Bus Interface block performs simple address decoding to select a peripheral and multiplexing of the returned signals from the various peripheral blocks. The base addresses used for the decode operation are defined in Table. Note that access to the MMU configuration registers are handled by the MMU Control Block rather than the CPU Subsystem Bus Interface block. The CPU Subsystem Bus Interface block operation is described by the following pseudocode:

masked_cpu_adr = cpu_adr[17:12] case (masked_cpu_adr) when TIM_base[17:12] cpu_tim_sel = peri_access_en  // The peri_access_en signal will have the peri_mmu_data = tim_cpu_data // timing required for block selects peri_mmu_rdy = tim_cpu_rdy peri_mmu_berr = tim_cpu_berr all_other_selects = 0  // Shorthand to ensure other cpu_block_sel signals // remain deasserted when LSS_base[17:12] cpu_lss_sel = peri_access_en peri_mmu_data = lss_cpu_data peri_mmu_rdy = lss_cpu_rdy peri_mmu_berr = lss_cpu_berr all_other_selects = 0 when GPIO_base[17:12] cpu_gpio_sel = peri_access_en peri_mmu_data = gpio_cpu_data peri_mmu_rdy = gpio_cpu_rdy peri_mmu_berr = gpio_cpu_berr all_other_selects = 0 when SCB_base[17:12] cpu_scb_sel = peri_access_en peri_mmu_data = scb_cpu_data peri_mmu_rdy = scb_cpu_rdy peri_mmu_berr = scb_cpu_berr all_other_selects = 0 when ICU_base[17:12] cpu_icu_sel = peri_access_en peri_mmu_data = icu_cpu_data peri_mmu_rdy = icu_cpu_rdy peri_mmu_berr = icu_cpu_berr all_other_selects = 0 when CPR_base[17:12] cpu_cpr_sel = peri_access_en peri_mmu_data = cpr_cpu_data peri_mmu_rdy = cpr_cpu_rdy peri_mmu_berr = cpr_cpu_berr all_other_selects = 0 when ROM_base[17:12] cpu_rom_sel = peri_access_en peri_mmu_data = rom_cpu_data peri_mmu_rdy = rom_cpu_rdy peri_mmu_berr = rom_cpu_berr all_other_selects = 0 when PSS_base[17:12] cpu_pss_sel = peri_access_en peri_mmu_data = pss_cpu_data peri_mmu_rdy = pss_cpu_rdy peri_mmu_berr = pss_cpu_berr all_other_selects = 0 when DIU_base[17:12] cpu_diu_sel = peri_access_en peri_mmu_data = diu_cpu_data peri_mmu_rdy = diu_cpu_rdy peri_mmu_berr = diu_cpu_berr all_other_selects = 0 when PCU_base[17:12] cpu_pcu_sel = peri_access_en peri_mmu_data = pcu_cpu_data peri_mmu_rdy = pcu_cpu_rdy peri_mmu_berr = pcu_cpu_berr all_other_selects = 0 when others all_block_selects = 0 peri_mmu_data = 0x00000000 peri_mmu_rdy = 0 peri_mmu_berr = 1 end case 11.6.6.3 MMU Control Block

The MMU Control Block determines whether every CPU access is a valid access. No more than one cycle is to be consumed in determining the validity of an access and all accesses must terminate with the assertion of either mmu_cpu_rdy or mmu_cpu_berr. To safeguard against stalling the CPU a simple bus timeout mechanism will be supported.

TABLE 24 MMU Control Block I/Os Port name Pins I/O Description Global SoPEC signals prst_n 1 In Global reset. Synchronous to pclk, active low. pclk 1 In Global clock Toplevel/Common MMU Control Block signals cpu_adr[21:2] 22 Out Address bus for both DRAM and peripheral access. cpu_acode[1:0] 2 Out CPU access code signals (cpu_mmu_acode) retimed to meet the CPU Subsystem Bus timing requirements dram_access_en 1 Out DRAM Access Enable signal. Indicates that the current CPU access is a valid DRAM access. MMU Control Block to LEON AHB bridge signals cpu_mmu_adr[31:0] 32 In CPU core address bus. cpu_dataout[31:0] 32 In Toplevel CPU data bus mmu_cpu_data[31:0] 32 Out Data bus to the CPU core. Carries the data for all CPU read operations cpu_rwn 1 In Toplevel CPU Read/notWrite signal. cpu_mmu_acode[1:0] 2 In CPU access code signals mmu_cpu_rdy 1 Out Ready signal to the CPU core. Indicates the completion of all valid CPU accesses. mmu_cpu_berr 1 Out Bus Error signal to the CPU core. This signal is asserted to terminate an invalid access. cpu_start_access 1 In Start Access signal from the LEON AHB bridge indicating the start of a data transfer and that the cpu_adr, cpu_dataout, cpu_rwn and cpu_acode signals are all valid. This signal is only asserted during the first cycle of an access. cpu_iack 1 In Interrupt Acknowledge signal from the CPU. This signal is only asserted during an interrupt acknowledge cycle. cpu_ben[1:0] 2 In Byte enable signals indicating which bytes of the 32- bit bus are being accessed. MMU Control Block to CPU Subsystem Bus Interface signals cpu_adr[17:12] 8 Out Toplevel CPU Address bus. Only bits 17-12 are required to decode the peripherals address space peri_access_en 1 Out Enable Access signal. A peripheral access cannot be initiated unless it has been enabled by the MMU Control Unit peri_mmu_data[31:0] 32 In Data bus from the selected peripheral peri_mmu_rdy 1 In Data Ready signal. Indicates the data on the peri_mmu_data bus is valid for a read cycle or that the data was successfully written to the peripheral for a write cycle. peri_mmu_berr 1 In Bus Error signal. Indicates a bus error has occurred in accessing the selected peripheral Description:

The MMU Control Block is responsible for the MMU's core functionality, namely determining whether or not an access to any part of the address map is valid. An access is considered valid if it is to a mapped area of the address space and if the CPU is running in the appropriate mode for that address space. Furthermore the MMU control block must correctly handle the special cases that are: an interrupt acknowledge cycle, a reset exception vector fetch, an access that crosses a 256-bit DRAM word boundary and a bus timeout condition. The following pseudocode shows the logic required to implement the MMU Control Block functionality. It does not deal with the timing relationships of the various signals—it is the designer's responsibility to ensure that these relationships are correct and comply with the different bus protocols. For simplicity the pseudocode is split up into numbered sections so that the functionality may be seen more easily.

It is important to note that the style used for the pseudocode will differ from the actual coding style used in the RTL implementation. The pseudocode is only intended to capture the required functionality, to clearly show the criteria that need to be tested rather than to describe how the implementation should be performed. In particular the different comparisons of the address used to determine which part of the memory map, which DRAM region (if applicable) and the permission checking should all be performed in parallel (with results ORed together where appropriate) rather than sequentially as the pseudocode implies.

PS0 Description: This first segment of code defines a number of constants and variables that are used elsewhere in this description. Most signals have been defined in the I/O descriptions of the MMU sub-blocks that precede this section of the document. The post_reset_state variable is used later (in section PS4) to determine if we should trap a null pointer access.

PS0: const UnusedBottom = 0x002AC000 const DRAMTop = 0x4027FFFF const UserDataSpace = b01 const UserProgramSpace = b00 const SupervisorDataSpace = b11 const SupervisorProgramSpace = b10 const ResetExceptionCycles = 0x2 cpu_adr_peri_masked[5:0] = cpu_mmu_adr[17:12] cpu_adr_dram_masked[16:0] = cpu_mmu_adr & 0x003FFFE0 if (prst_n == 0) then // Initialise everything cpu_adr = cpu_mmu_adr[21:2] peri_access_en = 0 dram_access_en = 0 mmu_cpu_data = peri_mmu_data mmu_cpu_rdy = 0 mmu_cpu_berr = 0 post_reset_state = TRUE access_initiated = FALSE cpu_access_cnt = 0 // The following is used to determine if we are coming out of reset for the purposes of // reset exception vector redirection. There may be a convenient signal in the CPU core // that we could use instead of this. if ((cpu_start_access == 1) AND (cpu_access_cnt < ResetExceptionCycles) AND (clock_tick == TRUE)) then cpu_access_cnt = cpu_access_cnt +1 else post_reset_state = FALSE

PS1 Description: This section is at the top of the hierarchy that determines the validity of an access. The address is tested to see which macro-region (i.e. Unused, CPU Subsystem or DRAM) it falls into or whether the reset exception vector is being accessed.

PS1: if (cpu_mmu_adr >= UnusedBottom) then // The access is to an invalid area of the address space. See section PS2 elsif ((cpu_mmu_adr > DRAMTop) AND (cpu_mmu_adr < UnusedBottom)) then // We are in the CPU Subsystem/PEP Subsystem address space. See section PS3 // Only remaining possibility is an access to DRAM address space // First we need to intercept the special case for the reset exception vector elsif (cpu_mmu_adr < 0x00000010) then // The reset exception is being accessed. See section PS4 elsif ((cpu_adr_dram_masked >= Region0Bottom) AND (cpu_adr_dram_masked <= Region0Top) ) then // We are in Region0. See section PS5 elsif ((cpu_adr_dram_masked >= RegionNBottom) AND (cpu_adr_dram_masked <= RegionNTop) ) then // we are in RegionN // Repeat the Region0 (i.e. section PS5) logic for each of Region1 to Region7 else // We could end up here if there were gaps in the DRAM regions peri_access_en = 0 dram_access_en = 0 mmu_cpu_berr = 1 // we have an unknown access error, most likely due to hitting mmu_cpu_rdy = 0 // a gap in the DRAM regions // Only thing remaining is to implement a bus timeout function. This is done in PS6 end

PS2 Description: Accesses to the large unused area of the address space are trapped by this section. No bus transactions are initiated and the mmu_cpu_berr signal is asserted.

PS2: elsif (cpu_mmu_adr >= UnusedBottom) then peri_access_en = 0 // The access is to an invalid area of the address space dram_access_en = 0 mmu_cpu_berr = 1 mmu_cpu_rdy = 0

PS3 Description: This section deals with accesses to CPU Subsystem peripherals, including the MMU itself. If the MMU registers are being accessed then no external bus transactions are required. Access to the MMU registers is only permitted if the CPU is making a data access from supervisor mode, otherwise a bus error is asserted and the access terminated. For non-MMU accesses then transactions occur over the CPU Subsystem Bus and each peripheral is responsible for determining whether or not the CPU is in the correct mode (based on the cpu_acode signals) to be permitted access to its registers. Note that all of the PEP registers are accessed via the PCU which is on the CPU Subsystem Bus.

PS3: elsif ((cpu_mmu_adr > DRAMTop) AND (cpu_mmu_adr < UnusedBottom)) then // We are in the CPU Subsystem/PEP Subsystem address space cpu_adr = cpu_mmu_adr[21:2] if (cpu_adr_peri_masked == MMU_base)  then // access is to local registers peri_access_en = 0 dram_access_en = 0 if (cpu_acode == SupervisorDataSpace) then for (i=0; i<26; i++) { if ((i == cpu_mmu_adr[6:2]) then // selects the addressed register if (cpu_rwn == 1) then mmu_cpu_data[16:0] = MMUReg[i] // MMUReg[i] is one of the mmu_cpu_rdy = 1 // registers in Table mmu_cpu_berr = 0 else // write cycle MMUReg[i] = cpu_dataout[16:0] mmu_cpu_rdy = 1 mmu_cpu_berr = 0 else // there is no register mapped to this address mmu_cpu_berr = 1 // do we really want a bus_error here as registers mmu_cpu_rdy = 0 // are just mirrored in other blocks else // we have an access violation mmu_cpu_berr = 1 mmu_cpu_rdy = 0 else // access is to something else on the CPU Subsystem Bus peri_access_en = 1 dram_access_en = 0 mmu_cpu_data = peri_mmu_data mmu_cpu_rdy = peri_mmu_rdy mmu_cpu_berr = peri_mmu_berr

PS4 Description: The only correct accesses to the locations beneath 0x00000010 are fetches of the reset trap handling routine and these should be the first accesses after reset. Here we trap all other accesses to these locations regardless of the CPU mode. The most likely cause of such an access will be the use of a null pointer in the program executing on the CPU.

PS4: elsif (cpu_mmu_adr < 0x00000010) then if (post_reset_state == TRUE)) then cpu adr = cpu mmu adr[21:2] peri_access_en = 1 dram_access_en = 0 mmu_cpu_data = peri_mmu_data mmu_cpu_rdy = peri_mmu_rdy mmu_cpu_berr = peri_mmu_berr else // we have a problem (almost certainly a null pointer) peri_access_en = 0 dram_access_en = 0 mmu_cpu_berr = 1 mmu_cpu_rdy = 0

ps5 Description: This large section of pseudocode simply checks whether the access is within the bounds of DRAM Region0 and if so whether or not the access is of a type permitted by the Region0Control register. If the access is permitted then a DRAM access is initiated. If the access is not of a type permitted by the Region0Control register then the access is terminated with a bus error.

PS5: elsif ((cpu_adr_dram_masked >= Region0Bottom) AND (cpu_adr_dram_masked <= Region0Top) ) then // we are in Region0 cpu_adr = cpu_mmu_adr[21:2] if (cpu_rwn == 1) then if ((cpu_acode == SupervisorProgramSpace AND Region0Control[2] == 1)) OR (cpu_acode == UserProgramSpace AND Region0Control[5] == 1)) then //  this is a valid instruction fetch from Region0 //  The dram_cpu_data bus goes directly to the LEON // AHB bridge which also handles the hready generation peri_access_en = 0 dram_access_en = 1 mmu_cpu_berr = 0 elsif ((cpu_acode == SupervisorDataSpace AND Region0Control[0] == 1)  OR (cpu_acode == UserDataSpace AND Region0Control[3] == 1)) then  // this is a valid read access from Region0 peri_access_en = 0 dram_access_en = 1 mmu_cpu_berr = 0 else  // we have an access violation peri_access_en = 0 dram_access_en = 0 mmu_cpu_berr = 1 mmu_cpu_rdy = 0 else // it is a write access if ((cpu_acode == SupervisorDataSpace AND Region0Control[1] == 1) OR (cpu_acode == UserDataSpace AND Region0Control[4] == 1)) then // this is a valid write access to Region0 peri_access_en = 0 dram_access_en = 1 mmu_cpu_berr = 0 else  // we have an access violation peri_access_en = 0 dram_access_en = 0 mmu_cpu_berr = 1 mmu_cpu_rdy = 0

PS6 Description: This final section of pseudocode deals with the special case of a bus timeout. This occurs when an access has been initiated but has not completed before the BusTimeout number of pclk cycles. While access to both DRAM and CPU/PEP Subsystem registers will take a variable number of cycles (due to DRAM traffic, PCU command execution or the different timing required to access registers in imported IP) each access should complete before a timeout occurs. Therefore it should not be possible to stall the CPU by locking either the CPU Subsystem or DIU buses. However given the fatal effect such a stall would have it is considered prudent to implement bus timeout detection.

PS6: // Only thing remaining is to implement a bus timeout function. if ((cpu_start_access == 1) then access_initiated = TRUE timeout_countdown = BusTimeout if ((mmu_cpu_rdy == 1 ) OR (mmu_cpu_berr ==1 )) then access_initiated = FALSE peri_access_en = 0 dram_access_en = 0 if ((clock_tick == TRUE) AND (access_initiated == TRUE) AND (BusTimeout != 0)) if (timeout_countdown > 0) then timeout_countdown −− else // timeout has occurred peri_access_en = 0 // abort the access dram_access_en = 0 mmu_cpu_berr = 1 mmu_cpu_rdy = 0 11.7 LEON Caches

The version of LEON implemented on SoPEC features 1 kB of ICache and 1 kB of DCache. Both caches are direct mapped and feature 8 word lines so their data RAMs are arranged as 32×256-bit and their tag RAMs as 32×30-bit (itag) or 32×32-bit (dtag). Like most of the rest of the LEON code used on SoPEC the cache controllers are taken from the leon2-1.0.7 release. The LEON cache controllers and cache RAMs have been modified to ensure that an entire 256-bit line is refilled at a time to make maximum use out of the memory bandwidth offered by the embedded DRAM organization (DRAM lines are also 256-bit). The data cache controller has also been modified to ensure that user mode code cannot access the DCache contents unless it is authorised to do so. A block diagram of the LEON CPU core as implemented on SoPEC is shown in FIG. 23 below.

In this diagram dotted lines are used to indicate hierarchy and red items represent signals or wrappers added as part of the SoPEC modifications. LEON makes heavy use of VHDL records and the records used in the CPU core are described in Table 25. Unless otherwise stated the records are defined in the iface.vhd file (part of the LEON release) and this should be consulted for a complete breakdown of the record elements.

TABLE 25 Relevant LEON records Record Name Description rfi Register File Input record. Contains address, datain and control signals for the register file. rfo Register File Output record. Contains the data out of the dual read port register file. ici Instruction Cache In record. Contains program counters from different stages of the pipeline and various control signals ico Instruction Cache Out record. Contains the fetched instruction data and various control signals. This record is also sent to the DCache (i.e. icol) so that diagnostic accesses (e.g. lda/sta) can be serviced. dci Data Cache In record. Contains address and data buses from different stages of the pipeline (execute & memory) and various control signals dco Data Cache Out record. Contains the data retrieved from either memory or the caches and various control signals. This record is also sent to the ICache (i.e. dcol) so that diagnostic accesses (e.g. lda/sta) can be serviced. iui Integer Unit In record. This record contains the interrupt request level and a record for use with LEONs Debug Support Unit (DSU) iuo Integer Unit Out record. This record contains the acknowledged interrupt request level with control signals and a record for use with LEONs Debug Support Unit (DSU) mcii Memory to Cache Icache In record. Contains the address of an Icache miss and various control signals mcio Memory to Cache Icache Out record. Contains the returned data from memory and various control signals mcdi Memory to Cache Dcache In record. Contains the address and data of a Dcache miss or write and various control signals mcdo Memory to Cache Dcache Out record. Contains the returned data from memory and various control signals ahbi AHB In record. This is the input record for an AHB master and contains the data bus and AHB control signals. The destination for the signals in this record is the AHB controller. This record is defined in the amba.vhd file ahbo AHB Out record. This is the output record for an AHB master and contains the address and data buses and AHB control signals. The AHB controller drives the signals in this record. This record is defined in the amba.vhd file ahbsi AHB Slave In record. This is the input record for an AHB slave and contains the address and data buses and AHB control signals. It is used by the DCache to facilitate cache snooping (this feature is not enabled in SoPEC). This record is defined in the amba.vhd file crami Cache RAM In record. This record is composed of records of records which contain the address, data and tag entries with associated control signals for both the ICache RAM and DCache RAM cramo Cache RAM Out record. This record is composed of records of records which contain the data and tag entries with associated control signals for both the ICache RAM and DCache RAM iline_rdy Control signal from the ICache controller to the instruction cache memory. This signal is active (high) when a full 256- bit line (on dram_cpu_data) is to be written to cache memory. dline_rdy Control signal from the DCache controller to the data cache memory. This signal is active (high) when a full 256-bit line (on dram_cpu_data) is to be written to cache memory. dram_cpu_data 256-bit data bus from the embedded DRAM 11.7.1 Cache Controllers

The LEON cache module consists of three components: the ICache controller (icache.vhd), the DCache controller (dcache.vhd) and the AHB bridge (acache.vhd) which translates all cache misses into memory requests on the AHB bus.

In order to enable full line refill operation a few changes had to be made to the cache controllers. The ICache controller was modified to ensure that whenever a location in the cache was updated (i.e. the cache was enabled and was being refilled from DRAM) all locations on that cache line had their valid bits set to reflect the fact that the full line was updated. The iline_rdy signal is asserted by the ICache controller when this happens and this informs the cache wrappers to update all locations in the idata RAM for that line.

A similar change was made to the DCache controller except that the entire line was only updated following a read miss and that existing write through operation was preserved. The DCache controller uses the dline_rdy signal to instruct the cache wrapper to update all locations in the ddata RAM for a line. An additional modification was also made to ensure that a double-word load instruction from a non-cached location would only result in one read access to the DIU i.e. the second read would be serviced by the data cache. Note that if the DCache is turned off then a double-word load instruction will cause two DIU read accesses to occur even though they will both be to the same 256-bit DRAM line.

The DCache controller was further modified to ensure that user mode code cannot access cached data to which it does not have permission (as determined by the relevant RegionNControl register settings at the time the cache line was loaded). This required an extra 2 bits of tag information to record the user read and write permissions for each cache line. These user access permissions can be updated in the same manner as the other tag fields (i.e. address and valid bits) namely by line refill, STA instruction or cache flush. The user access permission bits are checked every time user code attempts to access the data cache and if the permissions of the access do not agree with the permissions returned from the tag RAM then a cache miss occurs. As the MMU evaluates the access permissions for every cache miss it will generate the appropriate exception for the forced cache miss caused by the errant user code. In the case of a prohibited read access the trap will be immediate while a prohibited write access will result in a deferred trap. The deferred trap results from the fact that the prohibited write is committed to a write buffer in the DCache controller and program execution continues until the prohibited write is detected by the MMU which may be several cycles later. Because the errant write was treated as a write miss by the DCache controller (as it did not match the stored user access permissions) the cache contents were not updated and so remain coherent with the DRAM contents (which do not get updated because the MMU intercepted the prohibited write). Supervisor mode code is not subject to such checks and so has free access to the contents of the data cache.

In addition to AHB bridging, the ACache component also performs arbitration between ICache and DCache misses when simultaneous misses occur (the DCache always wins) and implements the Cache Control Register (CCR). The leon2-1.0.7 release is inconsistent in how it handles cacheability: For instruction fetches the cacheability (i.e. is the access to an area of memory that is cacheable) is determined by the ICache controller while the ACache determines whether or not a data access is cacheable. To further complicate matters the DCache controller does determine if an access resulting from a cache snoop by another AHB master is cacheable (Note that the SoPEC ASIC does not implement cache snooping as it has no need to do so). This inconsistency has been cleaned up in more recent LEON releases but is preserved here to minimise the number of changes to the LEON RTL. The cache controllers were modified to ensure that only DRAM accesses (as defined by the SoPEC memory map) are cached.

The only functionality removed as a result of the modifications was support for burst fills of the ICache. When enabled burst fills would refill an ICache line from the location where a miss occurred up to the end of the line. As the entire line is now refilled at once (when executing from DRAM) this functionality is no longer required. Furthermore more substantial modifications to the ICache controller would be needed if we wished to preserve this function without adversely affecting full line refills. The CCR was therefore modified to ensure that the instruction burst fetch bit (bit16) was tied low and could not be written to.

11.7.1.1 LEON Cache Control Register

The CCR controls the operation of both the I and D caches. Note that the bitfields used on the SoPEC implementation of this register are based on the LEON v1.0.7 implementation and some bits have their values tied off. See section 4 of the LEON manual for a description of the LEON cache controllers.

TABLE 26 LEON Cache Control Register Field Name bit(s) Description ICS 1:0 Instruction cache state: 00 - disabled 01 - frozen 10 - disabled 11 - enabled Reserved 13:6  Reserved. Reads as 0. DCS 3:2 Data cache state: 00 - disabled 01 - frozen 10 - disabled 11 - enabled IF 4 ICache freeze on interrupt 0 - Do not freeze the ICache contents on taking an interrupt 1 - Freeze the ICache contents on taking an interrupt DF 5 DCache freeze on interrupt 0 - Do not freeze the DCache contents on taking an interrupt 1 - Freeze the DCache contents on taking an interrupt Reserved 13:6  Reserved. Reads as 0. DP 14 Data cache flush pending. 0 - No DCache flush in progress 1 - DCache flush in progress This bit is ReadOnly. IP 15 Instruction cache flush pending. 0 - No ICache flush in progress 1 - ICache flush in progress This bit is ReadOnly. IB 16 Instruction burst fetch enable. This bit is tied low on SoPEC because it would interfere with the operation of the cache wrappers. Burst refill functionality is automatically provided in SoPEC by the cache wrappers. Reserved 20:17 Reserved. Reads as 0. FI 21 Flush instruction cache. Writing a 1 this bit will flush the ICache. Reads as 0. FD 22 Flush data cache. Writing a 1 this bit will flush the DCache. Reads as 0. DS 23 Data cache snoop enable. This bit is tied low in SoPEC as there is no requirement to snoop the data cache. Reserved 31:24 Reserved. Reads as 0. 11.7.2 Cache Wrappers

The cache RAMs used in the leon2-1.0.7 release needed to be modified to support full line refills and the correct IBM macros also needed to be instantiated. Although they are described as RAMs throughout this document (for consistency), register arrays are actually used to implement the cache RAMs. This is because IBM SRAMs were not available in suitable configurations (offered configurations were too big) to implement either the tag or data cache RAMs. Both instruction and data tag RAMs are implemented using dual port (1 Read & 1 Write) register arrays and the clocked write-through versions of the register arrays were used as they most closely approximate the single port SRAM LEON expects to see.

11.7.2.1 Cache Tag RAM Wrappers

The itag and dtag RAMs differ only in their width—the itag is a 32×30 array while the dtag is a 32×32 array with the extra 2 bits being used to record the user access permissions for each line. When read using a LDA instruction both tags return 32-bit words. The tag fields are described in Table 27 and Table 28 below. Using the IBM naming conventions the register arrays used for the tag RAMs are called RA032X30D2P2W1R1M3 for the itag and RA032X32D2P2W1R1M3 for the dtag. The ibm_syncram wrapper used for the tag RAMs is a simple affair that just maps the wrapper ports on to the appropriate ports of the IBM register array and ensures the output data has the correct timing by registering it. The tag RAMs do not require any special modifications to handle full line refills.

TABLE 27 LEON Instruction Cache Tag Field Name bit(s) Description Valid 7:0 Each valid bit indicates whether or not the corresponding word of the cache line contains valid data Reserved 9:8 Reserved - these bits do not exist in the itag RAM. Reads as 0. Address 31:10 The tag address of the cache line

TABLE 28 LEON Data Cache Tag Field Name bit(s) Description Valid 7:0 Each valid bit indicates whether or not the corresponding word of the cache line contains valid data URP 8 User read permission. 0 - User mode reads will force a refill of this line 1 - User mode code can read from this cache line. UWP 9 User write permission. 0 - User mode writes will not be written to the cache 1 - User mode code can write to this cache line. Address 31:10 The tag address of the cache line 11.7.2.2 Cache Data RAM Wrappers

The cache data RAM contains the actual cached data and nothing else. Both the instruction and data cache data RAMs are implemented using 8 32×32-bit register arrays and some additional logic to support full line refills. Using the IBM naming conventions the register arrays used for the tag RAMs are called RA032X32D2P2W1R1M3. The ibm_cdram_wrap wrapper used for the tag RAMs is shown in FIG. 24 below.

To the cache controllers the cache data RAM wrapper looks like a 256×32 single port SRAM (which is what they expect to see) with an input to indicate when a full line refill is taking place (the line_rdy signal). Internally the 8-bit address bus is split into a 5-bit lineaddress, which selects one of the 32 256-bit cache lines, and a 3-bit wordaddress which selects one of the 8 32-bit words on the cache line. Thus each of the 8 32×32 register arrays contains one 32-bit word of each cache line. When a full line is being refilled (indicated by both the line_rdy and write signals being high) every register array is written to with the appropriate 32 bits from the linedatain bus which contains the 256-bit line returned by the DIU after a cache miss. When just one word of the cache line is to be written (indicated by the write signal being high while the line_rdy is low) then the wordaddress is used to enable the write signal to the selected register array only—all other write enable signals are kept low. The data cache controller handles byte and half-word write by means of a read-modify-write operation so writes to the cache data RAM are always 32-bit.

The wordaddress is also used to select the correct 32-bit word from the cache line to return to the LEON integer unit.

11.8 Realtime Debug Unit (RDU)

The RDU facilitates the observation of the contents of most of the CPU addressable registers in the SoPEC device in addition to some pseudo-registers in realtime. The contents of pseudo-registers, i.e. registers that are collections of otherwise unobservable signals and that do not affect the functionality of a circuit, are defined in each block as required. Many blocks do not have pseudo-registers and some blocks (e.g. ROM, PSS) do not make debug information available to the RDU as it would be of little value in realtime debug.

Each block that supports realtime debug observation features a DebugSelect register that controls a local mux to determine which register is output on the block's data bus (i.e. block_cpu_data). One small drawback with reusing the blocks data bus is that the debug data cannot be present on the same bus during a CPU read from the block. An accompanying active high block_cpu_debug_valid signal is used to indicate when the data bus contains valid debug data and when the bus is being used by the CPU. There is no arbitration for the bus as the CPU will always have access when required. A block diagram of the RDU is shown in FIG. 25.

TABLE 29 RDU I/Os Port name Pins I/O Description diu_cpu_data 32 In Read data bus from the DIU block cpr_cpu_data 32 In Read data bus from the CPR block gpio_cpu_data 32 In Read data bus from the GPIO block icu_cpu_data 32 In Read data bus from the ICU block lss_cpu_data 32 In Read data bus from the LSS block pcu_cpu_debug_data 32 In Read data bus from the PCU block scb_cpu_data 32 In Read data bus from the SCB block tim_cpu_data 32 In Read data bus from the TIM block diu_cpu_debug_valid 1 In Signal indicating the data on the diu_cpu_data bus is valid debug data. tim_cpu_debug_valid 1 In Signal indicating the data on the tim_cpu_data bus is valid debug data. scb_cpu_debug_valid 1 In Signal indicating the data on the scb_cpu_data bus is valid debug data. pcu_cpu_debug_valid 1 In Signal indicating the data on the pcu_cpu_data bus is valid debug data. lss_cpu_debug_valid 1 In Signal indicating the data on the lss_cpu_data bus is valid debug data. icu_cpu_debug_valid 1 In Signal indicating the data on the icu_cpu_data bus is valid debug data. gpio_cpu_debug_valid 1 In Signal indicating the data on the gpio_cpu_data bus is valid debug data. cpr_cpu_debug_valid 1 In Signal indicating the data on the cpr_cpu_data bus is valid debug data. debug_data_out 32 Out Output debug data to be muxed on to the PHI/GPIO/other pins debug_data_valid 1 Out Debug valid signal indicating the validity of the data on debug_data_out. This signal is used in all debug configurations debug_cntrl 33 Out Control signal for each debug data line indicating whether or not the debug data should be selected by the pin mux

As there are no spare pins that can be used to output the debug data to an external capture device some of the existing I/Os will have a debug multiplexer placed in front of them to allow them be used as debug pins. Furthermore not every pin that has a debug mux will always be available to carry the debug data as they may be engaged in their primary purpose e.g. as a GPIO pin. The RDU therefore outputs a debug_cntrl signal with each debug data bit to indicate whether the mux associated with each debug pin should select the debug data or the normal data for the pin. The DebugPinSel1 and DebugPinSel2 registers are used to determine which of the 33 potential debug pins are enabled for debug at any particular time.

As it may not always be possible to output a full 32-bit debug word every cycle the RDU supports the outputting of an n-bit sub-word every cycle to the enabled debug pins. Each debug test would then need to be re-run a number of times with a different portion of the debug word being output on the n-bit sub-word each time. The data from each run should then be correlated to create a full 32-bit (or whatever size is needed) debug word for every cycle. The debug_data_valid and pclk_out signals will accompany every sub-word to allow the data to be sampled correctly. The pclk_out signal is sourced close to its output pad rather than in the RDU to minimise the skew between the rising edge of the debug data signals (which should be registered close to their output pads) and the rising edge of pclk_out.

As multiple debug runs will be needed to obtain a complete set of debug data the n-bit sub-word will need to contain a different bit pattern for each run. For maximum flexibility each debug pin has an associated DebugDataSrc register that allows any of the 32 bits of the debug data word to be output on that particular debug data pin. The debug data pin must be enabled for debug operation by having its corresponding bit in the DebugPinSel registers set for the selected debug data bit to appear on the pin.

The size of the sub-word is determined by the number of enabled debug pins which is controlled by the DebugPinSel registers. Note that the debug_data_valid signal is always output. Furthermore debug_cntrl[0] (which is configured by DebugPinSel1) controls the mux for both the debug_data_valid and pclk_out signals as both of these must be enabled for any debug operation. The mapping of debug_data_out[n] signals onto individual pins will take place outside the RDU.

This mapping is described in Table 30 below.

TABLE 30 DebugPinSel mapping bit # Pin DebugPinSel1 phi_frclk. The debug_data_valid signal will appear on this pin when enabled. Enabling this pin also automatically enables the phi_readl pin which will output the pclk_out signal DebugPinSel2(0-31) gpio[0 . . . 31]

TABLE 31 RDU Configuration Registers Address offset from MMU_base Register #bits Reset Description 0x80 DebugSrc 4 0x00 Denotes which block is supplying the debug data. The encoding of this block is given below. 0 - MMU 1 - TIM 2 - LSS 3 - GPIO 4 - SCB 5 - ICU 6 - CPR 7 - DIU 8 - PCU 0x84 DebugPinSel1 1 0x0 Determines whether the phi_frclk and phi_readl pins are used for debug output. 1 - Pin outputs debug data 0 - Normal pin function 0x88 DebugPinSel2 32 0x0000_0000 Determines whether a pin is used for debug data output. 1 - Pin outputs debug data 0 - Normal pin function 0x8C to DebugDataSrc[31:0] 32 × 5 0x00 Selects which bit of the 32-bit debug data 0x108 word will be output on debug_data_out[N] 11.9 Interrupt Operation

The interrupt controller unit (see chapter 14) generates an interrupt request by driving interrupt request lines with the appropriate interrupt level. LEON supports 15 levels of interrupt with level 15 as the highest level (the SPARC architecture manual [36] states that level 15 is non-maskable but we have the freedom to mask this if desired). The CPU will begin processing an interrupt exception when execution of the current instruction has completed and it will only do so if the interrupt level is higher than the current processor priority. If a second interrupt request arrives with the same level as an executing interrupt service routine then the exception will not be processed until the executing routine has completed.

When an interrupt trap occurs the LEON hardware will place the program counters (PC and nPC) into two local registers. The interrupt handler routine is expected, as a minimum, to place the PSR register in another local register to ensure that the LEON can correctly return to its pre-interrupt state. The 4-bit interrupt level (irl) is also written to the trap type (tt) field of the TBR (Trap Base Register) by hardware. The TBR then contains the vector of the trap handler routine the processor will then jump. The TBA (Trap Base Address) field of the TBR must have a valid value before any interrupt processing can occur so it should be configured at an early stage.

Interrupt pre-emption is supported while ET (Enable Traps) bit of the PSR is set. This bit is cleared during the initial trap processing. In initial simulations the ET bit was observed to be cleared for up to 30 cycles. This causes significant additional interrupt latency in the worst case where a higher priority interrupt arrives just as a lower priority one is taken.

The interrupt acknowledge cycles shown in FIG. 26 below are derived from simulations of the LEON processor. The SoPEC toplevel interrupt signals used in this diagram map directly to the LEON interrupt signals in the iui and iuo records. An interrupt is asserted by driving its (encoded) level on the icu_cpu_ilevel[3:0] signals (which map to iui.irl[3:0]). The LEON core responds to this, with variable timing, by reflecting the level of the taken interrupt on the cpu_icu_ilevel[3:0] signals (mapped to iuo.irl[3:0]) and asserting the acknowledge signal cpu_iack (iuo.intack). The interrupt controller then removes the interrupt level one cycle after it has seen the level been acknowledged by the core. If there is another pending interrupt (of lower priority) then this should be driven on icu_cpu_ilevel[3:0] and the CPU will take that interrupt (the level 9 interrupt in the example below) once it has finished processing the higher priority interrupt. The cpu_icu_ilevel[3:0] signals always reflect the level of the last taken interrupt, even when the CPU has finished processing all interrupts.

11.10 Boot Operation

See section 17.2 for a description of the SoPEC boot operation.

11.11 Software Debug

Software debug mechanisms are discussed in the “SoPEC Software Debug” document [15].

12 Serial Communications Block (SCB)

12.1 Overview

The Serial Communications Block (SCB) handles the movement of all data between the SoPEC and the host device (e.g. PC) and between master and slave SoPEC devices. The main components of the SCB are a Full-Speed (FS) USB Device Core, a FS USB Host Core, a Inter-SoPEC Interface (ISI), a DMA manager, the SCB Map and associated control logic. The need for these components and the various types of communication they provide is evident in a multi-SoPEC printer configuration.

12.1.1 Multi-SoPEC Systems

While single SoPEC systems are expected to form the majority of SoPEC systems the SoPEC device must also support its use in multi-SoPEC systems such as that shown in FIG. 27. A SoPEC may be assigned any one of a number of identities in a multi-SoPEC system. A SoPEC may be one or more of a PrintMaster, a LineSyncMaster, an ISIMaster, a StorageSoPEC or an ISISlave SoPEC.

12.1.1.1 ISIMaster Device

The ISIMaster is the only device that controls the common ISI lines (see FIG. 30) and typically interfaces directly with the host. In most systems the ISIMaster will simply be the SoPEC connected to the USB bus. Future systems, however, may employ an ISI-Bridge chip to interface between the host and the ISI bus and in such systems the ISI-Bridge chip will be the ISIMaster. There can only be one ISIMaster on an ISI bus.

Systems with multiple SoPECs may have more than one host connection, for example there could be two SoPECs communicating with the external host over their FS USB links (this would of course require two USB cables to be connected), but still only one ISIMaster.

While it is not expected to be required, it is possible for a device to hand over its role as the ISIMaster to another device on the ISI i.e. the ISIMaster is not necessarily fixed.

12.1.1.2 PrintMaster Device

The PrintMaster device is responsible for coordinating all aspects of the print operation. This includes starting the print operation in all printing SoPECs and communicating status back to the external host. When the ISIMaster is a SoPEC device it is also likely to be the PrintMaster as well. There may only be one PrintMaster in a system and it is most likely to be a SoPEC device.

12.1.1.3 LineSyncMaster Device

The LineSyncMaster device generates the lsync pulse that all SoPECs in the system must synchronize their line outputs with. Any SoPEC in the system could act as a LineSyncMaster although the PrintMaster is probably the most likely candidate. It is possible that the LineSyncMaster may not be a SoPEC device at all—it could, for example, come from some OEM motor control circuitry. There may only be one LineSyncMaster in a system.

12.1.1.4 Storage Device

For certain printer types it may be realistic to use one SoPEC as a storage device without using its print engine capability—that is to effectively use it as an ISI-attached DRAM. A storage SoPEC would receive data from the ISIMaster (most likely to be an ISI-Bridge chip) and then distribute it to the other SoPECs as required. No other type of data flow (e.g. ISISlave->storage SoPEC->ISISlave) would need to be supported in such a scenario. The SCB supports this functionality at no additional cost because the CPU handles the task of transferring outbound data from the embedded DRAM to the ISI transmit buffer. The CPU in a storage SoPEC will have almost nothing else to do.

12.1.1.5 ISISlave Device

Multi-SoPEC systems will contain one or more ISISlave SoPECs. An ISISlave SoPEC is primarily used to generate dot data for the printhead IC it is driving. An ISISlave will not transmit messages on the ISI without first receiving permission to do so, via a ping packet (see section 12.4.4.6), from the ISIMaster

12.1.1.6 ISI-Bridge Device

SoPEC is targeted at the low-cost small office/home office (SoHo) market. It may also be used in future systems that target different market segments which are likely to have a high speed interface capability. A future device, known as an ISI-Bridge chip, is envisaged which will feature both a high speed interface (such as High-Speed (HS) USB, Ethernet or IEEE1394) and one or more ISI interfaces. The use of multiple ISI buses would allow the construction of independent print systems within the one printer. The ISI-Bridge would be the ISIMaster for each of the ISI buses it interfaces to.

12.1.1.7 External Host

The external host is most likely (but is not required) to be, a PC. Any system that can act as a USB host or that can interface to an ISI-Bridge chip could be the external host. In particular, with the development of USB On-The-Go (USB OTG), it is possible that a number of USB OTG enabled products such as PDAs or digital cameras will be able to directly interface with a SoPEC printer.

12.1.1.8 External USB Device

The external USB device is most likely (but is not required) to be, a digital camera. Any system that can act as a USB device could be connected as an external USB device. This is to facilitate printing in the absence of a PC.

12.1.2 Types of Communication

12.1.2.1 Communications with External Host

The external host communicates directly with the ISIMaster in order to print pages. When the ISIMaster is a SoPEC, the communications channel is FS USB.

12.1.2.1.1 External Host to ISIMaster Communication

The external host will need to communicate the following information to the ISIMaster device:

-   -   Communications channel configuration and maintenance information     -   Most data destined for PrintMaster, ISISlave or storage SoPEC         devices. This data is simply relayed by the ISIMaster     -   Mapping of virtual communications channels, such as USB         endpoints, to ISI destination         12.1.2.1.2 ISIMaster to External Host Communication

The ISIMaster will need to communicate the following information to the external host:

-   -   Communications channel configuration and maintenance information     -   All data originating from the PrintMaster, ISISlave or storage         SoPEC devices and destined for the external host. This data is         simply relayed by the ISIMaster         12.1.2.1.3 External Host to PrintMaster Communication

The external host will need to communicate the following information to the PrintMaster device:

-   -   Program code for the PrintMaster     -   Compressed page data for the PrintMaster     -   Control messages to the PrintMaster     -   Tables and static data required for printing e.g. dead nozzle         tables, dither matrices etc.     -   Authenticatable messages to upgrade the printer's capabilities         12.1.2.1.4 PrintMaster to External Host Communication

The PrintMaster will need to communicate the following information to the external host:

-   -   Printer status information (i.e. authentication results, paper         empty/jammed etc.)     -   Dead nozzle information     -   Memory buffer status information     -   Power management status     -   Encrypted SoPEC_id for use in the generation of PRINTER_QA keys         during factory programming         12.1.2.1.5 External Host to ISISlave Communication

All communication between the external host and ISISlave SoPEC devices must be direct (via a dedicated connection between the external host and the ISISlave) or must take place via the ISIMaster. In the case of a SoPEC ISIMaster it is possible to configure each individual USB endpoint to act as a control channel to an ISISlave SoPEC if desired, although the endpoints will be more usually used to transport data. The external host will need to communicate the following information to ISISlave devices over the comms/lSI:

-   -   Program code for ISISlave SoPEC devices     -   Compressed page data for ISISlave SoPEC devices     -   Control messages to the ISISlave SoPEC (where a control channel         is supported)     -   Tables and static data required for printing e.g. dead nozzle         tables, dither matrices etc.     -   Authenticatable messages to upgrade the printer's capabilities         12.1.2.1.6 ISISlave to External Host Communication

All communication between the ISISlave SoPEC devices and the external host must take place via the ISIMaster. The ISISlave will need to communicate the following information to the external host over the comms/ISI:

-   -   Responses to the external host's control messages (where a         control channel is supported)     -   Dead nozzle information from the ISISlave SoPEC.     -   Encrypted SoPEC_id for use in the generation of PRINTER_QA keys         during factory programming         12.1.2.2 Communication with External USB Device         12.1.2.2.1 ISIMaster to External USB Device Communication     -   Communications channel configuration and maintenance         information.         12.1.2.2.2 External USB Device to ISIMaster Communication     -   Print data from a function on the external USB device.         12.1.2.3 Communication Over ISI         12.1.2.3.1 ISIMaster to PrintMaster Communication

The ISIMaster and PrintMaster will often be the same physical device. When they are different devices then the following information needs to be exchanged over the ISI:

-   -   All data from the external host destined for the PrintMaster         (see section 12.1.2.1.4).

This data is simply relayed by the ISIMaster

12.1.2.3.2 PrintMaster to ISIMaster Communication

The ISIMaster and PrintMaster will often be the same physical device. When they are different devices then the following information needs to be exchanged over the ISI:

-   -   All data from the PrintMaster destined for the external host         (see section 12.1.2.1.4).

This data is simply relayed by the ISIMaster

12.1.2.3.3 ISIMaster to ISISlave Communication

The ISIMaster may wish to communicate the following information to the ISISlaves:

-   -   All data (including program code such as ISIId enumeration)         originating from the external host and destined for the ISISlave         (see section 12.1.2.1.5). This data is simply relayed by the         ISIMaster     -   wake up from sleep mode         12.1.2.3.4 ISISlave to ISIMaster Communication

The ISISlave may wish to communicate the following information to the ISIMaster:

-   -   All data originating from the ISISlave and destined for the         external host (see section 12.1.2.1.6). This data is simply         relayed by the ISIMaster         12.1.2.3.5 PrintMaster to ISISlave Communication

When the PrintMaster is not the ISIMaster all ISI communication is done in response to ISI ping packets (see 12.4.4.6). When the PrintMaster is the ISIMaster then it will of course communicate directly with the ISISlaves. The PrintMaster SoPEC may wish to communicate the following information to the ISISlaves:

-   -   Ink status e.g. requests for dotCount data i.e. the number of         dots in each color fired by the printheads connected to the         ISISlaves     -   configuration of GPIO ports e.g. for clutch control and lid open         detect     -   power down command telling the ISISlave to enter sleep mode     -   ink cartridge fail information

This list is not complete and the time constraints associated with these requirements have yet to be determined.

In general the PrintMaster may need to be able to:

-   -   send messages to an ISISlave which will cause the ISISlave to         return the contents of ISISlave registers to the PrintMaster or     -   to program ISISlave registers with values sent by the         PrintMaster

This should be under the control of software running on the CPU which writes messages to the ISI/SCB interface.

12.1.2.3.6 ISISlave to PrintMaster Communication

ISISlaves may need to communicate the following information to the PrintMaster:

-   -   ink status e.g. dotCount data i.e. the number of dots in each         color fired by the printheads connected to the ISISlaves     -   band related information e.g. finished band interrupts     -   page related information i.e. buffer underrun, page finished         interrupts     -   MMU security violation interrupts     -   GPIO interrupts and status e.g. clutch control and lid open         detect     -   printhead temperature     -   printhead dead nozzle information from SoPEC printhead nozzle         tests     -   power management status

This list is not complete and the time constraints associated with these requirements have yet to be determined.

As the ISI is an insecure interface commands issued over the ISI should be of limited capability e.g. only limited register writes allowed. The software protocol needs to be constructed with this in mind. In general ISISlaves may need to return register or status messages to the PrintMaster or ISIMaster. They may also need to indicate to the PrintMaster or ISIMaster that a particular interrupt has occurred on the ISISlave. This should be under the control of software running on the CPU which writes messages to the ISI block.

12.1.2.3.7 ISISlave to ISISlave Communication

The amount of information that will need to be communicated between ISISlaves will vary considerably depending on the printer configuration. In some systems ISISlave devices will only need to exchange small amounts of control information with each other while in other systems (such as those employing a storage SoPEC or extra USB connection) large amounts of compressed page data may be moved between ISISlaves. ScenarIOs where ISISlave to ISISlave communication is required include: (a) when the PrintMaster is not the ISIMaster, (b) QA Chip ink usage protocols, (c) data transmission from data storage SoPECs, (d) when there are multiple external host connections supplying data to the printer.

12.1.3 SCB Block Diagram

The SCB consists of four main sub-blocks, as shown in the basic block diagram of FIG. 28.

12.1.4 Definitions of I/Os

The toplevel I/Os of the SCB are listed in Table 32. A more detailed description of their functionality will be given in the relevant sub-block sections.

TABLE 32 SCB I/O Port name s I/O Description Clocks and Resets prst_n 1 In System reset signal. Active low. Pclk 1 In System clock. usbclk 1 In 48 MHz clock for the USB device and host cores. The cores also require a 12 MHz clock, which will be generated locally by dividing the 48 MHz clock by 4. isi_cpr_reset_n 1 Out Signal from the ISI indicating that ISI activity has been detected while in sleep mode and so the chip should be reset. Active low. usbd_cpr_reset_n 1 Out Signal from the USB device that a USB reset has occurred. Active low. USB device IO transceiver signals usbd_ts 1 Out USB device IO transceiver (BUSB2_PM) driver three-state control. Active high enable. usbd_a 1 Out USB device IO transceiver (BUSB2_PM) driver data input. usbd_se0 1 Out USB device IO transceiver (BUSB2_PM) single-ended zero input. Active high. usbd_zp 1 In USB device IO transceiver (BUSB2_PM) D+ receiver output. usbd_zm 1 In USB device IO transceiver (BUSB2_PM) D− receiver output. usbd_z 1 In USB device IO transceiver (BUSB2_PM) differential receiver output. usbd_pull_up_en 1 Out USB device pull-up resistor enable. Switches power to the external pull-up resistor, connected to the D+ line that is required for device identification to the USB. Active high. usbd_vbus_sense 1 In USB device VBUS power sense. Used to detect power on VBUS. NOTE: The IBM Cu11 PADS are 3.3 V, VBUS is 5 V. An external volt- age conversion will be necessary, e.g. resistor divider network. Active high. USB host IO transceiver signals usbh_ts 1 Out USB host IO transceiver (BUSB2_PM) driver three-state control. Active high enable usbh_a 1 Out USB host IO transceiver (BUSB2_PM) driver data input. usbh_se0 1 Out USB host IO transceiver (BUSB2_PM) single- ended zero input. Active high. usbh_zp 1 In USB host IO transceiver (BUSB2_PM) D+ receiver output. usbh_zm 1 In USB host IO transceiver (BUSB2_PM) D− receiver output. usbh_z 1 In USB host IO transceiver (BUSB2_PM) differential receiver output. usbh_over_current 1 In USB host port power over current indicator. Active high. usbh_power_en 1 Out USB host VBUS power enable. Used for port power switching. Active high. CPU Interface cpu_adr[n:2] n − 1 In CPU address bus. cpu_dataout[31:0] 32 In Shared write data bus from the CPU scb_cpu_data[31:0] 32 Out Read data bus to the CPU cpu_rwn 1 In Common read/not-write signal from the CPU cpu_acode[1:0] 2 In CPU Access Code signals. These decode as follows: 00 - User program access 01 - User data access 10 - Supervisor program access 11 - Supervisor data access cpu_scb_sel 1 In Block select from the CPU. When cpu_scb_sel is high both cpu_adr and cpu_dataout are valid scb_cpu_rdy 1 Out Ready signal to the CPU. When scb_cpu_rdy is high it indicates the last cycle of the access. For a write cycle this means cpu_dataout has been registered by the SCB and for a read cycle this means the data on scb_cpu_data is valid. scb_cpu_berr 1 Out Bus error signal to the CPU indicating an invalid access. scb_cpu_debug_valid 1 Out Signal indicating that the data currently on scb_cpu_data is valid debug data Interrupt signals dma_icu_irq 1 Out DMA interrupt signal to the interrupt controller block. isi_icu_irq 1 Out ISI interrupt signal to the interrupt controller block. usb_icu_irq[1:0] 2 Out USB host and device interrupt signals to the ICU. Bit 0 - USB Host interrupt Bit 1 - USB Device interrupt DIU interface scb_diu_wadr[21:5] 17 Out Write address bus to the DIU scb_diu_data[63:0] 64 Out Data bus to the DIU. scb_diu_wreq 1 Out Write request to the DIU diu_scb_wack 1 In Acknowledge from the DIU that the write request was accepted. scb_diu_wvalid 1 Out Signal from the SCB to the DIU indicating that the data currently on the scb_diu_data[63:0] bus is valid scb_diu_wmask[7:0] 7 Out Byte aligned write mask. A “1” in a bit field of “scb_diu_wmask[7:0]” means that the corresponding byte will be written to DRAM. scb_diu_rreq 1 Out Read request to the DIU. scb_diu_radr[21:5] 17 Out Read address bus to the DIU diu_scb_rack 1 In Acknowledge from the DIU that the read request was accepted. diu_scb_rvalid 1 In Signal from the DIU to the SCB indicating that the data currently on the diu_data [63:0] bus is valid diu_data[63:0] 64 In Common DIU data bus. GPIO interface isi_gpio_dout[3:0] 4 Out ISI output data to GPIO pins isi_gpio_e[3:0] 4 Out ISI output enable to GPIO pins gpio_isi_din[3:0] 4 In Input data from GPIO pins to ISI 12.1.5 SCB Data Flow

A logical view of the SCB is shown in FIG. 29, depicting the transfer of data within the SCB.

12.2 USBD (USB Device Sub-Block)

12.2.1 Overview

The FS USB device controller core and associated SCB logic are referred to as the USB Device (USBD).

A SoPEC printer has FS USB device capability to facilitate communication between an external USB host and a SoPEC printer. The USBD is self-powered. It connects to an external USB host via a dedicated USB interface on the SoPEC printer, comprising a USB connector, the necessary discretes for USB signalling and the associated SoPEC ASIC I/Os.

The FS USB device core will be third party IP from Synopsys: TymeWare™ USB1.1 Device Controller (UDCVCI). Refer to the UDCVCI User Manual [20] for a description of the core.

The device core does not support LS USB operation. Control and bulk transfers are supported by the device. Interrupt transfers are not considered necessary because the required interrupt-type functionality can be achieved by sending query messages over the control channel on a scheduled basis. There is no requirement to support isochronous transfers.

The device core is configured to support 6 USB endpoints (EPs): the default control EP (EP0), 4 bulk OUT EPs (EP1, EP2, EP3, EP4) and 1 bulk IN EP (EP5). It should be noted that the direction of each EP is with respect to the USB host, i.e. IN refers to data transferred to the external host and OUT refers to data transferred from the external host. The 4 bulk OUT EPs will be used for the transfer of data from the external host to SoPEC, e.g. compressed page data, program data or control messages. Each bulk OUT EP can be mapped on to any target destination in a multi-SoPEC system, via the SCB Map configuration registers. The bulk IN EP is used for the transfer of data from SoPEC to the external host, e.g. a print image downloaded from a digital camera that requires processing on the external host system. Any feedback data will be returned to the external host on EP0, e.g. status information.

The device core does not provide internal buffering for any of its EPs (with the exception of the 8 byte setup data payload for control transfers). All EP buffers are provided in the SCB. Buffers will be grouped according to EP direction and associated packet destination. The SCB Map configuration registers contain a DestISIId and DestISISubId for each OUT EP, defining their EP mapping and therefore their packet destination. Refer to section Section 12.4 ISI (Inter SoPEC Interface Sub-block) for further details on ISIId and ISISubId. Refer to section Section 12.5 CTRL (Control Sub-block) for further details on the mapping of OUT EPs.

12.2.2 USBD Effective Bandwidth

The effective bandwidth between an external USB host and the printer will be influenced by:

-   -   Amount of activity from other devices that share the USB with         the printer.     -   Throughput of the device controller core.     -   EP buffering implementation.     -   Responsiveness of the external host system CPU in handling USB         interrupts.

To maximize bandwidth to the printer it is recommended that no other devices are active on the USB between the printer and the external host. If the printer is connected to a HS USB external host or hub it may limit the bandwidth available to other devices connected to the same hub but it would not significantly affect the bandwidth available to other devices upstream of the hub. The EP buffering should not limit the USB device core throughput, under normal operating conditions. Used in the recommended configuration, under ideal operating conditions, it is expected that an effective bandwidth of 8-9 Mbit/s will be achieved with bulk transfers between the external host and the printer.

12.2.3 IN EP Packet Buffer

The IN EP packet buffer stores packets originating from the LEON CPU that are destined for transmission over the USB to the external USB host. CPU writes to the buffer are 32 bits wide. USB device core reads from the buffer 32 bits wide.

128 bytes of local memory are required in total for EP0-IN and EP5-IN buffering. The IN EP buffer is a single, 2-port local memory instance, with a dedicated read port and a dedicated write port. Both ports are 32 bits wide. Each IN EP has a dedicated 64 byte packet location available in the memory array to buffer a single USB packet (maximum USB packet size is 64 bytes). Each individual 64 byte packet location is structured as 16×32 bit words and is read/written in a FIFO manner. When the device core reads a packet entry from the IN EP packet buffer, the buffer must retain the packet until the device core performs a status write, informing the SCB that the packet has been accepted by the external USB host and can be flushed. The CPU can therefore only write a single packet at a time to each IN EP. Any subsequent CPU write request to a buffer location containing a valid packet will be refused, until that packet has been successfully transmitted.

12.2.4 OUT EP Packet Buffer

The OUT EP packet buffer stores packets originating from the external USB host that are destined for transmission over DMAChannel0, DMAChannel1 or the ISI. The SCB control logic is responsible for routing the OUT EP packets from the OUT EP packet buffer to DMA or to the ISITx Buffer, based on the SCB Map configuration register settings. USB core writes to the buffer are 32 bits wide. DMA and ISI associated reads from the buffer are both 64 bits wide.

512 bytes of local memory are required in total for EP0-OUT, EP1-OUT, EP2-OUT, EP3-OUT and EP4-OUT buffering. The OUT EP packet buffer is a single, 2-port local memory instance, with a dedicated read port and a dedicated write port. Both ports are 64 bits wide. Byte enables are used for the 32 bit wide USB device core writes to the buffer. Each OUT EP can be mapped to DMAChannel0, DMAChannel1 or the ISI.

The OUT EP packet buffer is partitioned accordingly, resulting in three distinct packet FIFOs:

-   -   USBDDMA0FIFO, for USB packets destined for DMAChannel0 on the         local SoPEC.     -   USBDDMA1 FIFO, for USB packets destined for DMAChannel1 on the         local SoPEC.     -   USBDISIFIFO, for USB packets destined for transmission over the         ISI.         12.2.4.1 USBDDMAnFIFO

This description applies to USBDDMA0FIFO and USBDDMA1 FIFO, where ‘n’ represents the respective DMA channel, i.e. n=0 for USBDDMA0FIFO, n=1 for USBDDMA1 FIFO. USBDDMAnFIFO services any EPs mapped to DMAChanneln on the local SoPEC device. This implies that a packet originating from an EP with an associated ISIId that matches the local SoPEC ISIId and an ISISubId=n will be written to USBDDMAnFIFO, if there is space available for that packet.

USBDDMAnFIFO has a capacity of 2×64 byte packet entries, and can therefore buffer up to 2 USB packets. It can be considered as a 2 packet entry FIFO. Packets will be read from it in the same order in which they were written, i.e. the first packet written will be the first packet read and the second packet written will be the second packet read. Each individual 64 byte packet location is structured as 8×64 bit words and is read/written in a FIFO manner.

The USBDDMAnFIFO has a write granularity of 64 bytes, to allow for the maximum USB packet size. The USBDDMAnFIFO will have a read granularity of 32 bytes to allow for the DMA write access bursts of 4×64 bit words, i.e. the DMA Manager will read 32 byte chunks at a time from the USBDDMAnFIFO 64 byte packet entries, for transfer to the DIU.

It is conceivable that a packet which is not a multiple 32 bytes in size may be written to the USBDDMAnFIFO. When this event occurs, the DMA Manager will read the contents of the remaining address locations associated with the 32 byte chunk in the USBDDMAnFIFO, transferring the packet plus whatever data is present in those locations, resulting in a 32 byte packet (a burst of 4×64 bit words) transfer to the DIU.

The DMA channels should achieve an effective bandwidth of 160 Mbits/sec (1 bit/cycle) and should never become blocked, under normal operating conditions. As the USB bandwidth is considerably less, a 2 entry packet FIFO for each DMA channel should be sufficient.

12.2.4.2 USBDISIFIFO

USBDISIFIFO services any EPs mapped to ISI. This implies that a packet originating from an EP with an associated ISId that does not match the local SoPEC ISId will be written to USBDISIFIFO if there is space available for that packet.

USBDISIFIFO has a capacity of 4×64 byte packet entries, and can therefore buffer up to 4 USB packets. It can be considered as a 4 packet entry FIFO. Packets will be read from it in the same order in which they were written, i.e. the first packet written will be the first packet read and the second packet written will be the second packet read, etc. Each individual 64 byte packet location is structured as 8×64 bit words and is read/written in a FIFO manner.

The ISI long packet format will be used to transfer data across the ISI. Each ISI long packet data payload is 32 bytes. The USBDISIFIFO has a write granularity of 64 bytes, to allow for the maximum USB packet size. The USBDISIFIFO will have a read granularity of 32 bytes to allow for the ISI packet size, i.e. the SCB will read 32 byte chunks at a time from the USBDISIFIFO 64 byte packet entries, for transfer to the ISI.

It is conceivable that a packet which is not a multiple 32 bytes in size may be written to the USBDISIFIFO, either intentionally or due to a software error. A maskable interrupt per EP is provided to flag this event. There will be 2 options for dealing with this scenario on a per EP basis:

-   -   Discard the packet.     -   Read the contents of the remaining address locations associated         with the 32 byte chunk in the USBDISIFIFO, transferring the         irregular size packet plus whatever data is present in those         locations, resulting in a 32 byte packet transfer to the         ISITxBuffer.

The ISI should achieve an effective bandwidth of 100 Mbits/sec (4 wire configuration). It is possible to encounter a number of retries when transmitting an ISI packet and the LEON CPU will require access to the ISI transmit buffer. However, considering the relatively low bandwidth of the USB, a 4 packet entry FIFO should be sufficient.

12.2.5 Wake-Up From Sleep Mode

The SoPEC will be placed in sleep mode after a suspend command is received by the USB device core. The USB device core will continue to be powered and clocked in sleep mode. A USB reset, as opposed to a device resume, will be required to bring SoPEC out of its sleep state as the sleep state is hoped to be logically equivalent to the power down state.

The USB reset signal originating from the USB controller will be propagated to the CPR (as usb_cpr_reset_n) if the USBWakeupEnable bit of the WakeupEnable register (see Table) has been set. The USBWakeupEnable bit should therefore be set just prior to entering sleep mode. There is a scenario that would require SoPEC to initiate a USB remote wake-up (i.e. where SoPEC signals resume to the external USB host after being suspended by the external USB host). A digital camera (or other supported external USB device) could be connected to SoPEC via the internal SoPEC USB host controller core interface. There may be a need to transfer data from this external USB device, via SoPEC, to the external USB host system for processing. If the USB connecting the external host system and SoPEC was suspended, then SoPEC would need to initiate a USB remote wake-up.

12.2.6 Implementation

12.2.6.1 USBD Sub-Block Partition

-   -   Block diagram     -   Definition of I/Os     -   12.2.6.2 USB Device IP Core     -   12.2.6.3 PVCI Target     -   12.2.6.4 IN EP Buffer     -   12.2.6.5 OUT EP Buffer     -   12.3 USBH (USB Host Sub-Block)     -   12.3.1 Overview

The SoPEC USB Host Controller (HC) core, associated SCB logic and associated SoPEC ASIC I/Os are referred to as the USB Host (USBH).

A SoPEC printer has FS USB host capability, to facilitate communication between an external USB device and a SoPEC printer. The USBH connects to an external USB device via a dedicated USB interface on the SoPEC printer, comprising a USB connector, the necessary discretes for USB signalling and the associated SoPEC ASIC I/Os.

The FS USB HC core are third party IP from Synopsys: DesignWare^(R) USB1.1 OHCI Host Controller with PVCI (UHOSTC_PVCI). Refer to the UHOSTC_PVCI User Manual [18] for details of the core. Refer to the Open Host Controller Interface (OHCI) Specification Release [19] for details of OHCI operation.

The HC core supports Low-Speed (LS) USB devices, although compatible external USB devices are most likely to be FS devices. It is expected that communication between an external USB device and a SoPEC printer will be achieved with control and bulk transfers. However, isochronous and interrupt transfers are also supported by the HC core.

There will be 2 communication channels between the Host Controller Driver (HCD) software running on the LEON CPU and the HC core:

-   -   OHCI operational registers in the HC core. These registers are         control, status, list pointers and a pointer to the Host         Controller Communications Area (HCCA) in shared memory. A target         Peripheral Virtual Component Interface (PCVI) on the HC core         will provide LEON with direct read/write access to the         operational registers. Refer to the OHCI Specification for         details of these registers.     -   HCCA in SoPEC eDRAM. An initiator Peripheral Virtual Component         Interface (PCVI) on the HC core will provide the HC with DMA         read/write access to an address space in eDRAM. The HCD running         on LEON will have read/write access to the same address space.         Refer to the OHCI Specification for details of the HCCA.

The target PVCI interface is a 32 bit word aligned interface, with byte enables for write access. All read/write access to the target PVCI interface by the LEON CPU will be 32 bit word aligned. The byte enables will not be used, as all registers will be read and written as 32 bit words.

The initiator PVCI interface is a 32 bit word aligned interface with byte enables for write access. All DMA read/write accesses are 256 bit word aligned, in bursts of 4×64 bit words. As there is no guarantee that the read/write requests from the HC core will start at a 256 bit boundary or be 256 bits long, it is necessary to provide 8 byte enables for each of the 64 bit words in a write burst form the HC core to DMA. The signal scb_diu_wmask serves this purpose.

Configuration of the HC core will be performed by the HCD.

12.3.2 Read/Write Buffering

The HC core maximum burst size for a read/write access is 4×32 bit words. This implies that the minimum buffering requirements for the HC core will be a 1 entry deep address register and a 4 entry deep data register. It will be necessary to provide data and address mapping functionality to convert the 4×32 bit word HC core read/write bursts into 4×64 bit word DMA read/write bursts. This will meet the minimum buffering requirements.

12.3.3 USBH Effective Bandwidth

The effective bandwidth between an external USB device and a SoPEC printer will be influenced by:

-   -   Amount of activity from other devices that share the USB with         the external USB device.     -   Throughput of the HC core.     -   HC read/write buffering implementation.     -   Responsiveness of the LEON CPU in handling USB interrupts.

Effective bandwidth between an external USB device and a SoPEC printer is not an issue. The primary application of this connectivity is the download of a print image from a digital camera. Printing speed is not important for this type of print operation. However, to maximize bandwidth to the printer it is recommended that no other devices are active on the USB between the printer and the external USB device. The HC read/write buffering in the SCB should not limit the USB HC core throughput, under normal operating conditions.

Used in the recommended configuration, under ideal operating conditions, it is expected that an effective bandwidth of 8-9 Mbit/s will be achieved with bulk transfers between the external USB device and the SoPEC printer.

12.3.4 Implementation

12.3.5 USBH Sub-Block Partition

USBH Block Diagram

Definition of I/Os.

12.3.5.1 USB Host IP Core

12.3.5.2 PVCI Target

12.3.5.3 PVCI Initiator

12.3.5.4 Read/Write Buffer

12.4 ISI (Inter SoPEC Interface Sub-Block)

12.4.1 Overview

The ISI is utilised in all system configurations requiring more than one SoPEC. An example of such a system which requires four SoPECs for duplex A3 printing and an additional SoPEC used as a storage device is shown in FIG. 27.

The ISI performs much the same function between an ISISlave SoPEC and the ISIMaster as the USB connection performs between the ISIMaster and the external host. This includes the transfer of all program data, compressed page data and message (i.e. commands or status information) passing between the ISIMaster and the ISISlave SoPECs. The ISIMaster initiates all communication with the ISISlaves.

12.4.2 ISI Effective Bandwidth

The ISI will need to run at a speed that will allow error free transmission on the PCB while minimising the buffering and hardware requirements on SoPEC. While an ISI speed of 10 Mbit/s is adequate to match the effective FS USB bandwidth it would limit the system performance when a high-speed connection (e.g. USB2.0, IEEE1394) is used to attach the printer to the PC. Although they would require the use of an extra ISI-Bridge chip such systems are envisaged for more expensive printers (compared to the low-cost basic SoPEC powered printers that are initially being targeted) in the future.

An ISI line speed (i.e. the speed of each individual ISI wire) of 32 Mbit/s is therefore proposed as it will allow ISI data to be over-sampled 5 times (at a pclk frequency of 160 MHz). The total bandwidth of the ISI will depend on the number of pins used to implement the interface. The ISI protocol will work equally well if 2 or 4 pins are used for transmission/reception. The ISINumPins register is used to select between a 2 or 4 wire ISI, giving peak raw bandwidths of 64 Mbit/s and 128 Mbit/s respectively. Using either a 2 or 4 wire ISI solution would allow the movement of data in to and out of a storage SoPEC (as described in 12.1.1.4 above), which is the most bandwidth hungry ISI use, in a timely fashion.

The ISINumPins register is used to select between a 2 or 4 wire ISI. A 2 wire ISI is the default setting for ISINumPins and this may be changed to a 4 wire ISI after initial communication has been established between the ISIMaster and all ISISlaves. Software needs to ensure that the switch from 2 to 4 wires is handled in a controlled and coordinated fashion so that nothing is transmitted on the ISI during the switch over period.

The maximum effective bandwidth of a two wire ISI, after allowing for protocol overheads and bus turnaround times, is expected to be approx. 50 Mbit/s.

12.4.3 ISI Device Identification and Enumeration

The ISIMasterSel bit of the ISICntrl register (see section Table) determines whether a SoPEC is an ISIMaster (ISIMasterSel=1), or an ISISlave (ISIMasterSel=0).

SoPEC defaults to being an ISISlave (ISIMasterSel=0) after a power-on reset—i.e. it will not transmit data on the ISI without first receiving a ping. If a SoPEC's ISIMasterSel bit is changed to 1, then that SoPEC will become the ISIMaster, transmitting data without requiring a ping, and generating pings as appropriately programmed.

ISIMasterSel can be set to 1 explicitly by the CPU writing directly to the ISICntrl register. ISIMasterSel can also be automatically set to 1 when activity occurs on any of USB endpoints 2-4 and the AutoMasterEnable bit of the ISICntrl register is also 1 (the default reset condition). Note that if AutoMasterEnable is 0, then activity on USB endpoints 2-4 will not result in ISIMasterSel being set to 1. USB endpoints 2-4 are chosen for the automatic detection since the power-on-reset condition has USB endpoints 0 and 1 pointing to ISIId 0 (which matches the local SoPEC's ISIId after power-on reset). Thus any transmission on USB endpoints 2-4 indicate a desire to transmit on the ISI which would usually indicate ISIMaster status. The automatic setting of ISIMasterSel can be disabled by clearing AutoMasterEnable, thereby allowing the SoPEC to remain an ISISlave while still making use of the USB endpoints 2-4 as external destinations.

Thus the setting of a SoPEC being ISIMaster or ISISlave can be completely under software control, or can be completely automatic.

The ISIId is established by software downloaded over the ISI (in broadcast mode) which looks at the input levels on a number of GPIO pins to determine the ISIId. For any given printer that uses a multi-SoPEC configuration it is expected that there will always be enough free GPIO pins on the ISISlaves to support this enumeration mechanism.

12.4.4 ISI Protocol

The ISI is a serial interface utilizing a 2/4 wire half-duplex configuration such as the 2-wire system shown in FIG. 30 below. An ISIMaster must always be present and a variable number of ISISlaves may also be on the ISI bus. The ISI protocol supports up to 14 addressable slaves, however to simplify electrical issues the ISI drivers need only allow for 5-6 ISI devices on a particular ISI bus. The ISI bus enables broadcasting of data, ISIMaster to ISISlave communication, ISISlave to ISIMaster communication and ISISlave to ISISlave communication. Flow control, error detection and retransmission of errored packets is also supported. ISI transmission is asynchronous and a Start field is present in every transmitted packet to ensure synchronization for the duration of the packet.

To maximize the effective ISI bandwidth while minimising pin requirements a half-duplex interleaved transmission scheme is used. FIG. 31 below shows how a 16-bit word is transmitted from an ISIMaster to an ISISlave over a 2-wire ISI bus. Since data will be interleaved over the wires and a 4-wire ISI is also supported, all ISI packets should be a multiple of 4 bits.

All ISI transactions are initiated by the ISIMaster and every non-broadcast data packet needs to be acknowledged by the addressed recipient. An ISISlave may only transmit when it receives a ping packet (see section 12.4.4.6) addressed to it. To avoid bus contention all ISI devices must wait ISITurnAround bit-times (5 pclk cycles per bit) after detecting the end of a packet before transmitting a packet (assuming they are required to transmit). All non-transmitting ISI devices must tristate their Tx drivers to avoid line contention. The ISI protocol is defined to avoid devices driving out of order (e.g. when an ISISlave is no longer being addressed). As the ISI uses standard I/O pads there is no physical collision detection mechanism.

There are three types of ISI packet: a long packet (used for data transmission), a ping packet (used by the ISIMaster to prompt ISISlaves for packets) and a short packet (used to acknowledge receipt of a packet). All ISI packets are delineated by a Start and Stop fields and transmission is atomic i.e. an ISI packet may not be split or halted once transmission has started.

12.4.4.1 ISI Transactions

The different types of ISI transactions are outlined in FIG. 32 below. As described later all NAKs are inferred and ACKS are not addressed to any particular ISI device.

12.4.4.2 Start Field Description

The Start field serves two purposes: To allow the start of a packet be unambiguously identified and to allow the receiving device synchronise to the data stream. The symbol, or data value, used to identify a Start field must not legitimately occur in the ensuing packet. Bit stuffing is used to guarantee that the Start symbol will be unique in any valid (i.e. error free) packet. The ISI needs to see a valid Start symbol before packet reception can commence i.e. the receive logic constantly looks for a Start symbol in the incoming data and will reject all data until it sees a Start symbol. Furthermore if a Start symbol occurs (incorrectly) during a data packet it will be treated as the start of a new packet. In this case the partially received packet will be discarded.

The data value of the Start symbol should guarantee that an adequate number of transitions occur on the physical ISI lines to allow the receiving ISI device to determine the best sampling window for the transmitted data. The Start symbol should also be sufficiently long to ensure that the bit stuffing overhead is low but should still be short enough to reduce its own contribution to the packet overhead. A Start symbol of b01010101 is therefore used as it is an effective compromise between these constraints.

Each SoPEC in a multi-SoPEC system will derive its system clock from a unique (i.e. one per SoPEC) crystal. The system clocks of each device will drift relative to each other over any period of time. The system clocks are used for generation and sampling of the ISI data. Therefore the sampling window can drift and could result in incorrect data values being sampled at a later point in time. To overcome this problem the ISI receive circuitry tracks the sampling window against the incoming data to ensure that the data is sampled in the centre of the bit period.

12.4.4.3 Stop Field Description

A 1 bit-time Stop field of b1 per ISI line ensures that all ISI lines return to the high state before the next packet is transmitted. The stop field is driven on to each ISI line simultaneously, i.e. b11 for a 2-wire ISI and b1111 for a 4-wire ISI would be interleaved over the respective ISI lines. Each ISI line is driven high for 1 bit-time. This is necessary because the first bit of the Start field is b0.

12.4.4.4 Bit Stuffing

This involves the insertion of bits into the bitstream at the transmitting SoPEC to avoid certain data patterns. The receiving SoPEC will strip these inserted bits from the bitstream.

Bit-stuffing will be performed when the Start symbol appears at a location other than the start field of any packet, i.e. when the bit pattern b0101010 occurs at the transmitter, a 0 will be inserted to escape the Start symbol, resulting in the bit pattern b01010100. Conversely, when the bit pattern b0101010 occurs at the receiver, if the next bit is a ‘0’ it will be stripped, if it is a ‘1’ then a Start symbol is detected.

If the frequency variations in the quartz crystal were large enough, it is conceivable that the resultant frequency drift over a large number of consecutive 1s or 0s could cause the receiving SoPEC to loose synchronisation.⁶ The quartz crystal that will be used in SoPEC systems is rated for 32 MHz @ 100 ppm. In a multi-SoPEC system with a 32 MHz+100 ppm crystal and a 32 MHz-100 ppm crystal, it would take approximately 5000 pclk cycles to cause a drift of 1 pclk cycle. This means that we would only need to bit-stuff somewhere before 1000 ISI bits of consecutive 1s or consecutive 0s, to ensure adequate synchronization. As the maximum number of bits transmitted per ISI line in a packet is 145, it should not be necessary to perform bit-stuffing for consecutive 1s or 0s. We may wish to constrain the spec of xtalin and also xtalin for the ISI-Bridge chip to ensure the ISI cannot drift out of sync during packet reception. ⁶Current max packet size˜=290 bits=145 bits per ISI line (on a 2 wire ISI)=725 160 MHz cycles. Thus the pclks in the two communicating ISI devices should not drift by more than one cycle in 725 i.e. 1379 ppm. Careful analysis of the crystal, PLL and oscillator specs and the sync detection circuit is needed here to ensure our solution is robust.

Note that any violation of bit stuffing will result in the RxFrameErrorSticky status bit being set and the incoming packet will be treated as an errored packet.

12.4.4.5 ISI Long Packet

The format of a long ISI packet is shown in FIG. 33 below. Data may only be transferred between ISI devices using a long packet as both the short and ping packets have no payload field. Except in the case of a broadcast packet, the receiving ISI device will always reply to a long packet with an explicit ACK (if no error is detected in the received packet) or will not reply at all (e.g. an error is detected in the received packet), leaving the transmitter to infer a NAK. As with all ISI packets the bitstream of a long packet is transmitted with its lsb (the leftmost bit in FIG. 33) first. Note that the total length (in bits) of an ISI long packet differs slightly between a 2 and 4-wire ISI system due to the different number of bits required for the Start and Stop fields.

All long packets begin with the Start field as described earlier. The PktDesc field is described in

TABLE 33 PktDesc field description Bit Description 0:1 00 - Long packet 01 - Reserved 10 - Ping packet 11 - Reserved 2 Sequence bit value. Only valid for long packets. See section 12.4.4.9 for a description of sequence bit operation

Any ISI device in the system may transmit a long packet but only the ISIMaster may initiate an ISI transaction using a long packet. An ISISlave may only send a long packet in reply to a ping message from the ISIMaster. A long packet from an ISISlave may be addressed to any ISI device in the system.

The Address field is straightforward and complies with the ISI naming convention described in section 12.5.

The payload field is exactly what is in the transmit buffer of the transmitting ISI device and gets copied into the receive buffer of the addressed ISI device(s). When present the payload field is always 256 bits.

To ensure strong error detection a 16-bit CRC is appended.

12.4.4.6 ISI Ping Packet

The ISI ping packet is used to allow ISISlaves to transmit on the ISI bus. As can be seen from FIG. 34 below the ping packet can be viewed as a special case of the long packet. In other words it is a long packet without any payload. Therefore the PktDesc field is the same as a long packet PktDesc, with the exception of the sequence bit, which is not valid for a ping packet. Both the ISISubId and the sequence bit are fixed at 1 for all ping packets. These values were chosen to maximize the hamming distance from an ACK symbol and to minimize the likelihood of bit stuffing. The ISISubId is unused in ping packets because the ISIMaster is addressing the ISI device rather than one of the DMA channels in the device. The ISISlave may address any ISIId.ISISubId in response if it wishes. The ISISlave will respond to a ping packet with either an explicit ACK (if it has nothing to send), an inferred NAK (if it detected an error in the ping packet) or a long packet (containing the data it wishes to send). Note that inferred NAKS do not result in the retransmission of a ping packet. This is because the ping packet will be retransmitted on a predetermined schedule (see 12.4.4.11 for more details).

An ISISlave should never respond to a ping message to the broadcast ISIId as this must have been sent in error. An ISI ping packet will never be sent in response to any packet and may only originate from an ISIMaster.

12.4.4.7 ISI Short Packet

The ISI short packet is only 17 bits long, including the Start and Stop fields. A value of b11101011 is proposed for the ACK symbol. As a 16-bit CRC is inappropriate for such a short packet it is not used. In fact there is only one valid value for a short ACK packet as the Start, ACK and Stop symbols all have fixed values. Short packets are only used for acknowledgements (i.e. explicit ACKs). The format of a short ISI packet is shown in FIG. 35 below. The ACK value is chosen to ensure that no bit stuffing is required in the packet and to minimize its hamming distance from ping and long ISI packets.

12.4.4.8 Error Detection and Retransmission

The 16-bit CRC will provide a high degree of error detection and the probability of transmission errors occurring is very low as the transmission channel (i.e. PCB traces) will have a low inherent bit error rate. The number of undetected errors should therefore be minute.

The HDLC standard CRC-16 (i.e. G(x)=x¹⁶+x¹²+x⁵+1) is to be used for this calculation, which is to be performed serially. It is calculated over the entire packet (excluding the Start and Stop fields). A simple retransmission mechanism frees the CPU from getting involved in error recovery for most errors because the probability of a transmission error occurring more than once in succession is very, very low in normal circumstances.

After each non-short ISI packet is transmitted the transmitting device will open a reply window. The size of the reply window will be ISIShortReplyWin bit times when a short packet is expected in reply, i.e. the size of a short packet, allowing for worst case bit stuffing, bus turnarounds and timing differences. The size of the reply window will be ISILongReplyWin bit times when a long packet is expected in reply, i.e. this will be the max size of a long packet, allowing for worst case bit stuffing, bus turnarounds and timing differences. In both cases if an ACK is received the window will close and another packet can be transmitted but if an ACK is not received then the full length of the window must be waited out.

As no reply should be sent to a broadcast packet, no reply window should be required however all other long packets open a reply window in anticipation of an ACK. While the desire is to minimize the time between broadcast transmissions the simplest solution should be employed. This would imply the same size reply window as other long packets.

When a packet has been received without any errors the receiving ISI device must transmit its acknowledge packet (which may be either a long or short packet) before the reply window closes. When detected errors do occur the receiving ISI device will not send any response. The transmitting ISI device interprets this lack of response as a NAK indicating that errors were detected in the transmitted packet or that the receiving device was unable to receive the packet for some reason (e.g. its buffers are full). If a long packet was transmitted the transmitting ISI device will keep the transmitted packet in its transmit buffer for retransmission. If the transmitting device is the ISIMaster it will retransmit the packet immediately while if the transmitting device is an ISISlave it will retransmit the packet in response to the next ping it receives from the ISIMaster.

The transmitting ISI device will continue retransmitting the packet when it receives a NAK until it either receives an ACK or the number of retransmission attempts equals the value of the NumRetries register. If the transmission was unsuccessful then the transmitting device sets the TxErrorSticky bit in its ISIIntStatus register. The receiving device also sets the RxErrorSticky bit in its ISIIntStatus register whenever it detects a CRC error in an incoming packet and is not required to take any further action, as it is up to the transmitting device to detect and rectify the problem. The NumRetries registers in all ISI devices should be set to the same value for consistent operation. Note that successful transmission or reception of ping packets do not affect retransmission operation.

Note that a transmit error will cause the ISI to stop transmitting. CPU intervention will be required to resolve the source of the problem and to restart the ISI transmit operation. Receive errors however do not affect receive operation and they are collected to facilitate problem debug and to monitor the quality of the ISI physical channel. Transmit or receive errors should be extremely rare and their occurrence will most likely indicate a serious problem.

Note that broadcast packets are never acknowledged to avoid contention on the common ISI lines. If an ISISlave detects an error in a broadcast packet it should use the message passing mechanism described earlier to alert the ISIMaster to the error if it so wishes.

12.4.4.9 Sequence Bit Operation

To ensure that communication between transmitting and receiving ISI devices is correctly ordered a sequence bit is included in every long packet to keep both devices in step with each other. The sequence bit field is a constant for short or ping packets as they are not used for data transmission. In addition to the transmitted sequence bit all ISI devices keep two local sequence bits, one for each ISISubId. Furthermore each ISI device maintains a transmit sequence bit for each ISIId and ISISubId it is in communication with. For packets sourced from the external host (via USB) the transmit sequence bit is contained in the relevant USBEPnDest register while for packets sourced from the CPU the transmit sequence bit is contained in the CPUISITxBuffCntrl register. The sequence bits for received packets are stored in ISISubId0Seq and ISISubId1Seq registers. All ISI devices will initialize their sequence bits to 0 after reset. It is-the responsibility of software to ensure that the sequence bits of the transmitting and receiving ISI devices are correctly initialized each time a new source is selected for any ISIId.ISISubId channel.

Sequence bits are ignored by the receiving ISI device for broadcast packets. However the broadcasting ISI device is free to toggle the sequence in the broadcast packets since they will not affect operation. The SCB will do this for all USB source data so that there is no special treatment for the sequence bit of a broadcast packet in the transmitting device. CPU sourced broadcasts will have sequence bits toggled at the discretion of the program code.

Each SoPEC may also ignore the sequence bit on either of its ISISubId channels by setting the appropriate bit in the ISISubIdSeqMask register. The sequence bit should be ignored for ISISubId channels that will carry data that can originate from more than one source and is self ordering e.g. control messages.

A receiving ISI device will toggle its sequence bit addressed by the ISISubId only when the receiver is able to accept data and receives an error-free data packet addressed to it. The transmitting ISI device will toggle its sequence bit for that ISIId.ISISubId channel only when it receives a valid ACK handshake from the addressed ISI device.

FIG. 36 shows the transmission of two long packets with the sequence bit in both the transmitting and receiving devices toggling from 0 to 1 and back to 0 again. The toggling operation will continue in this manner in every subsequent transmission until an error condition is encountered.

When the receiving ISI device detects an error in the transmitted long packet or is unable to accept the packet (because of full buffers for example) it will not return any packet and it will not toggle its local sequence bit. An example of this is depicted in FIG. 37. The absence of any response prompts the transmitting device to retransmit the original (seq=0) packet. This time the packet is received without any errors (or buffer space may have been freed) so the receiving ISI device toggles its local sequence bit and responds with an ACK. The transmitting device then toggles its local sequence bit to a 1 upon correct receipt of the ACK.

However it is also possible for the ACK packet from the receiving ISI device to be corrupted and this scenario is shown in FIG. 38. In this case the receiving device toggles its local sequence bit to 1 when the long packet is received without error and replies with an ACK to the transmitting device. The transmitting device does not receive the ACK correctly and so does not change its local sequence bit. It then retransmits the seq=0 long packet. When the receiving device finds that there is a mismatch between the transmitted sequence bit and the expected (local) sequence bit is discards the long packet and replies with an ACK. When the transmitting ISI device correctly receives the ACK it updates its local sequence bit to a 1, thus restoring synchronization. Note that when the ISISubIdSeqMask bit for the addressed ISISubId is set then the retransmitted packet is not discarded and so a duplicate packet will be received. The data contained in the packet should be self-ordering and so the software handling these packets (most likely control messages) is expected to deal with this eventuality.

12.4.4.10 Flow Control

The ISI also supports flow control by treating it in exactly the same manner as an error in the received packet. Because the SCB enjoys greater guaranteed bandwidth to DRAM than both the ISI and USB can supply flow control should not be required during normal operation. Any blockage on a DMA channel will soon result in the NumRetries value being exceeded and transmission from that SoPEC being halted. If a SoPEC NAKS a packet because its RxBuffer is full it will flag an overflow condition. This condition can potentially cause a CPU interrupt, if the corresponding interrupt is enabled. The RxOverflowSticky bit of its ISIIntStatus register reflects this condition. Because flow control is treated in the same manner as an error the transmitting ISI device will not be able to differentiate a flow control condition from an error in the transmitted packet.

12.4.4.11 Auto-Ping Operation

While the CPU of the ISIMaster could send a ping packet by writing the appropriate header to the CPUISITxBuffCntrl register it is expected that all ping packets will be generated in the ISI itself. The use of automatically generated ping packets ensures that ISISlaves will be given access to the ISI bus with a programmable minimum guaranteed frequency in addition to whenever it would otherwise be idle. Five registers facilitate the automatic generation of ping messages within the ISI: PingSchedule0, PingSchedule1, PingSchedule2, ISITotalPeriod and ISILocalPeriod. Auto-pinging will be enabled if any bit of any of the PingScheduleN registers is set and disabled if all PingScheduleN registers are 0x0000.

Each bit of the 15-bit PingScheduleN register corresponds to an ISIId that is used in the Address field of the ping packet and a 1 in the bit position indicates that a ping packet is to be generated for that ISIId. A 0 in any bit position will ensure that no ping packet is generated for that ISIId. As ISISlaves may differ in their bandwidth requirement (particularly if a storage SoPEC is present) three different PingSchedule registers are used to allow an ISISlave receive up to three times the number of pings as another active ISISlave. When the ISIMaster is not sending long packets (sourced from either the CPU or USB in the case of a SoPEC ISIMaster) ISI ping packets will be transmitted according to the pattern given by the three PingScheduleN registers. The ISI will start with the lsb of PingSchedule0 register and work its way from lsb through msb of each of the PingScheduleN registers. When the msb of PingSchedule2 is reached the ISI returns to the lsb of PingSchedule0 and continues to cycle through each bit position of each PingScheduleN register. The ISI has more than enough time to work out the destination of the next ping packet while a ping or long packet is being transmitted.

With the addition of auto-ping operation we now have three potential sources of packets in an ISIMaster SoPEC: USB, CPU and auto-ping. Arbitration between the CPU and USB for access to the ISI is handled outside the ISI. To ensure that local packets get priority whenever possible and that ping packets can have some guaranteed access to the ISI we use two 4-bit counters whose reload value is contained in the ISITotalPeriod and ISILocalPeriod registers. As we saw in section 12.4.4.1 every ISI transaction is initiated by the ISIMaster transmitting either a long packet or a ping packet. The ISITotalPeriod counter is decremented for every ISI transaction (i.e. either long or ping) when its value is non-zero. The ISILocalPeriod counter is decremented for every local packet that is transmitted. Neither counter is decremented by a retransmitted packet. If the ISITotalPeriod counter is zero then ping packets will not change its value from zero. Both the ISITotalPeriod and ISILocalPeriod counters are reloaded by the next local packet transmit request after the ISITotalPeriod counter has reached zero and this local packet has priority over pings.

The amount of guaranteed ISI bandwidth allocated to both local and ping packets is determined by the values of the ISITotalPeriod and ISILocalPeriod registers. Local packets will always be given priority when the ISILocalPeriod counter is non-zero. Ping packets will be given priority when the ISILocalPeriod counter is zero and the ISITotalPeriod counter is still non-zero.

Note that ping packets are very likely to get more than their guaranteed bandwidth as they will be transmitted whenever the ISI bus would otherwise be idle (i.e. no pending local packets). In particular when the ISITotalPeriod counter is zero it will not be reloaded until another local packet is pending and so ping packets transmitted when the ISITotalPeriod counter is zero will be in addition to the guaranteed bandwidth. Local packets on the other hand will never get more than their guaranteed bandwidth because each local packet transmitted decrements both counters and will cause the counters to be reloaded when the ISITotalPeriod counter is zero. The difference between the values of the ISITotalPeriod and ISILocalPeriod registers determines the number of automatically generated ping packets that are guaranteed to be transmitted every ISITotalPeriod number of ISI transactions. If the ISITotalPeriod and ISILocalPeriod values are the same then the local packets will always get priority and could totally exclude ping packets if the CPU always has packets to send.

For example if ISITotalPeriod=0xC; ISILocalPeriod=0x8; PingSchedule0=0x0E; PingSchedule1=0x0C and PingSchedule2=0x08 then four ping messages are guaranteed to be sent in every 12 ISI transactions. Furthermore ISIId3 will receive 3 times the number of ping packets as ISId1 and ISId2 will receive twice as many as ISId1. Thus over a period of 36 contended ISI transactions (allowing for two full rotations through the three PingScheduleN registers) when local packets are always pending 24 local packets will be sent, ISId1 will receive 2 ping packets, ISId2 will receive 4 pings and ISId3 will receive 6 ping packets. If local traffic is less frequent then the ping frequency will automatically adjust upwards to consume all remaining ISI bandwidth.

12.4.5 Wake-Up From Sleep Mode

Either the PrintMaster SoPEC or the external host may place any of the ISISlave SoPECs in sleep mode prior to going into sleep mode itself. The ISISlave device should then ensure that its ISIWakeupEnable bit of the WakeupEnable register (see Table 34) is set prior to entering sleep mode. In an ISISlave device the ISI block will continue to receive power and clock during sleep mode so that it may monitor the gpio_isi_din lines for activity. When ISI activity is detected during sleep mode and the ISIWakeupEnable bit is set the ISI asserts the isi_cpr_reset_n signal. This will bring the rest of the chip out of sleep mode by means of a wakeup reset. See chapter 16 for more details of reset propagation.

12.4.6 Implementation

Although the ISI consists of either 2 or 4 ISI data lines over which a serial data stream is demultiplexed, each ISI line is treated as a separate serial link at the physical layer. This permits a certain amount of skew between the ISI lines that could not be tolerated if the lines were treated as a parallel bus. A lower Bit Error Rate (BER) can be achieved if the serial data recovery is performed separately on each serial link. FIG. 39 illustrates the ISI sub block partitioning.

12.4.6.1 ISI Sub-Block Partition

-   -   Definition of I/Os.

TABLE 34 ISI I/O Port name Pins I/O Description Clock and Reset isi_pclk 1 In ISI primary clock. isi_reset_n 1 In ISI reset. Active low. Asserting isi_reset_n will reset all ISI logic. Synchronous to isi_pclk. Configuration isi_go 1 In ISI GO. Active high. When GO is de-asserted, all ISI statemachines are reset to their idle states, all ISI output signals are de- asserted, but all ISI counters retain their values. When GO is asserted, all ISI counters are reset and all ISI statemachines and output signals will return to their normal mode of operation. isi_master_select 1 In ISI master select. Determines whether the SoPEC is an ISIMaster or not 1 = ISIMaster 0 = ISISlave isi_id[3:0] 4 In ISI ID for this device. isi_retries[3:0] 4 In ISI number of retries. Number of times a transmitting ISI device will attempt retransmission of a NAK'd packet before aborting the transmission and flagging an error. The value of this configuration signal should not be changed while there are valid packets in the Tx buffer. isi_ping_schedule0[14:0] 15 In ISI auto ping schedule #0. Denotes which ISIIds will be receive ping packets. Note that bit0 refers to ISIId0, bit1 to ISIId1 . . . bit14 to ISIId14. Setting a bit in this schedule will enable auto ping generation for the corresponding ISI ID. The ISI will start from the bit 0 of isi_ping_schedule0 and cycle through to bit 14, generating pings for each bit that is set. This operation will be performed in sequence from isi_ping_schedule0 through isi_ping_schedule2. isi_ping_schedule1[14:0] 15 In As per isi_ping_schedule0. isi_ping_schedule2[14:0] 15 In As per isi_ping_schedule0. isi_total_period[3:0] 4 In Reload value of the ISI Total Period Counter. isi_local_period[3:0] 4 In Reload value of the ISI Local Period Counter. isi_number_pins 1 In Number of active ISI data pins. Used to select how many serial data pins will be used to transmit and receive data. Should reflect the number of ISI device data pins that are in use. 1 = isi_data[3:0] active 0 = isi_data[1:0] active isi_turn_around[3:0] 4 In ISI bus turn around time in ISI clock cycles (32 MHz). isi_short_reply_win[4:0] 5 In ISI long packet reply window in ISI clock cycles (32 MHz). isi_long_reply_win[8:0] 9 In ISI long packet reply window in ISI clock cycles (32 MHz). isi_tx_enable 1 In ISI transmit enable. Active high. Enables ISI transmission of long or ping packets. ACKs may still be transmitted when this bit is 0. The value of this configuration signal should not be changed while there are valid packets in the Tx buffer. isi_rx_enable 1 In ISI receive enable. Active high. Enables ISI packet reception. Any activity on the ISI bus will be ignored when this signal is de-asserted. This signal should only be de-asserted if the ISI block is not required for use in the design. isi_bit_stuff_rate[3:0] 1 In ISI bit stuffing limit. Allows the bit stuffing counter value to be programmed. Is loaded into the 4 upper bits of the 7 bit wide bit stuffing counter. The lower bits are always loaded with b111, to prevent bit stuffing for less than 7 consecutive ones or zeroes. E.g. b000: stuff_count = b0000111: bit stuff after 7 consecutive 0/1 b111: stuff_count = b1111111: bit stuff after 127 consecutive 0/1 Serial Link Signals isi_ser_data_in[3:0] 4 In ISI Serial data inputs. Each bit corresponds to a separate serial link. isi_ser_data_out[3:0] 4 Out ISI Serial data outputs. Each bit corresponds to a separate serial link. isi_ser_data_en[3:0] 4 Out ISI Serial data driver enables. Active high. Each bit corresponds to a separate serial link. Tx Packet Buffer isi_tx_wr_en 1 In ISI Tx FIFO write enable. Active high. Asserting isi_tx_wr_en will write the 64 bit data on isi_tx_wr_data to the FIFO, providing that space is available in the FIFO. If isi_tx_wr_en remains asserted after the last entry in the current packet is written, the write operation will wrap around to the start of the next packet, providing that space is available for a second packet in the FIFO. isi_tx_wr_data[63:0] 64 In ISI Tx FIFO write data. isi_tx_ping 1 In ISI Tx FIFO ping packet select. Active high. Asserting isi_tx_ping will queue a ping packet for transmission, as opposed to a long packet. Although there is no data payload for a ping packet, a packet location in the FIFO is used as a ‘place holder’ for the ping packet. Any data written to the associated packet location in the FIFO will be discarded when the ping packet is transmitted. isi_tx_id[3:0] 5 In ISI Tx FIFO packet ID. ISI ID for each packet written to the FIFO. Registered when the last entry of the packet is written. isi_tx_sub_id 1 In ISI Tx FIFO packet sub ID. ISI sub ID for each packet written to the FIFO. Registered when the last entry of the packet is written. isi_tx_pkt_count[1:0] 2 Out ISI Tx FIFO packet count. Indicates the number of packets contained in the FIFO. The FIFO has a capacity of 2 × 256 bit packets. Range is b00−>b10. isi_tx_word_count[2:0] 3 Out ISI Tx FIFO current packet word count. Indicates the number of words contained in the current Tx packet location of the Tx FIFO. Each packet location has a capacity of 4 × 64 bit words. Range is b000−> b100. isi_tx_empty 1 Out ISI Tx FIFO empty. Active high. Indicates that no packets are present in the FIFO. isi_tx_full 1 Out ISI Tx FIFO full. Active high. Indicates that 2 packets are present in the FIFO, therefore no more packets can be transmitted. isi_tx_over_flow 1 Out ISI Tx FIFO over flow. Active high. Indicates that a write operation was performed on a full FIFO. The write operation will have no effect on the contents of the FIFO or the write pointer. isi_tx_error 1 Out ISI Tx FIFO error. Active high. Indicates that an error occurred while transmitting the packet currently at the head of the FIFO. This will happen if the number of transmission attempts exceeds isi_tx_retries. isi_tx_desc[2:0] 3 Out ISI Tx packet descriptor field. ISI packet descriptor field for the packet currently at the head of the FIFO. See Table for details. Only valid when isi_tx_empty = 0, i.e. when there is a valid packet in the FIFO. isi_tx_addr[4:0] 5 Out ISI Tx packet address field. ISI address field for the packet currently at the head of the FIFO. See Table for details. Only valid when isi_tx_empty = 0, i.e. when there is a valid packet in the FIFO. Rx Packet FIFO isi_rx_rd_en 1 In ISI Rx FIFO read enable. Active high. Asserting isi_rx_rd_en will drive isi_rx_rd_data with valid data, from the Rx packet at the head of the FIFO, providing that data is available in the FIFO. If isi_rx_rd_en remains asserted after the last entry is read from the current packet, the read operation will wrap around to the start of the next packet, providing that a second packet is available in the FIFO. isi_rx_rd_data[63:0] 64 Out ISI Rx FIFO read data. isi_rx_sub_id 1 Out ISI Rx packet sub ID. Indicates the ISI sub ID associated with the packet at the head of the Rx FIFO. isi_rx_pkt_count[1:0] 2 Out ISI Rx FIFO packet count. Indicates the number of packets contained in the FIFO. The FIFO has a capacity of 2 × 256 bit packets. Range is b00−>b10. isi_rx_word_count[2:0] 3 Out ISI Rx FIFO current packet word count. Indicates the number of words contained in the Rx packet location at the head of the FIFO. Each packet location has a capacity of 4 × 64 bit words. Range is b000−>b100. isi_rx_empty 1 Out ISI Rx FIFO empty. Active high. Indicates that no packets are present in the FIFO. isi_rx_full 1 Out ISI Rx FIFO full. Active high. Indicates that 2 packets are present in the FIFO, therefore no more packets can be received. isi_rx_over_flow 1 Out ISI Rx FIFO over flow. Active high. Indicates that a packet was addressed to the local ISI device, but the Rx FIFO was full, resulting in a NAK. isi_rx_under_run 1 Out ISI Rx FIFO under run. Active high. Indicates that a read operation was performed on an empty FIFO. The invalid read will return the contents of the memory location currently addressed by the FIFO read pointer and will have no effect on the read pointer. isi_rx_frame_error 1 Out ISI Rx framing error. Active high. Asserted by the ISI when a framing error is detected in the received packet, which can be caused by an incorrect Start or Stop field or by bit stuffing errors. The associated packet will be dropped. isi_rx_crc_error 1 Out ISI Rx CRC error. Active high. Asserted by the ISI when a CRC error is detected in an incoming packet. Other than dropping the errored packet ISI reception is unaffected by a CRC Error. 12.4.6.2 ISI Serial Interface Engine (isi_sie)

There are 4 instantiations of the isi_sie sub block in the ISI, 1 per ISI serial link. The isi_sie is responsible for Rx serial data sampling, Tx serial data output and bit stuffing.

Data is sampled based on a phase detection mechanism. The incoming ISI serial data stream is over sampled 5 times per ISI bit period. The phase of the incoming data is determined by detecting transitions in the ISI serial data stream, which indicates the ISI bit boundaries. An ISI bit boundary is defined as the sample phase at which a transition was detected.

The basic functional components of the isi_sie are detailed in FIG. 40. These components are simply a grouping of logical functionality and do not necessarily represent hierarchy in the design.

12.4.6.2.1 SIE Edge Detection and Data I/O

The basic structure of the data I/O and edge detection mechanism is detailed in FIG. 41.

NOTE: Serial data from the receiver in the pad MUST be synchronized to the isi_pclk domain with a 2 stage shift register external to the ISI, to reduce the risk of metastability. ser_data_out and ser_data_en should be registered externally to the ISI.

The Rx/Tx statemachine drives ser_data_en, stuff_(—)1_en and stuff_(—)0_en. The signals stuff_(—)1_en and stuff_(—)0_en cause a one or a zero to be driven on ser_data_out when they are asserted, otherwise fifo_rd_data is selected.

12.4.6.2.2 SIE Rx/Tx Statemachine

The Rx/Tx statemachine is responsible for the transmission of ISI Tx data and the sampling of ISI Rx data. Each ISI bit period is 5 isi_pclk cycles in duration.

The Tx cycle of the Rx/Tx statemachine is illustrated in FIG. 42. It generates each ISI bit that is transmitted. States tx0->tx4 represent each of the 5 isi_pclk phases that constitute a Tx ISI bit period. ser_data_en controls the tristate enable for the ISI line driver in the bidirectional pad, as shown in FIG. 41. rx_tx_cycle is asserted during both Rx and Tx states to indicate an active Rx or Tx cycle. It is primarily used to enable bit stuffing.

NOTE: All statemachine signals are assumed to be ‘0’ unless otherwise stated.

The Tx cycle for Tx bit stuffing when the Rx/Tx statemachine inserts a ‘0’ into the bitstream can be seen in FIG. 43.

NOTE: All statemachine signals are assumed to be ‘0’ unless otherwise stated

The Tx cycle for Tx bit stuffing when the RxTx statemachine inserts a ‘1’ into the bitstream can be seen in FIG. 44.

NOTE: All statemachine signals are assumed to be ‘0’ unless otherwise stated

The tx* and stuff* states are detailed separately for clarity. They could be easily combined when coding the statemachine, however it would be better for verification and debugging if they were kept separate.

The Rx cycle of the ISI Rx/Tx statemachine is detailed in FIG. 45. The Rx cycle of the Rx/Tx Statemachine, samples each ISI bit that is received. States rx0->rx4 represent each of the 5 isi_pclk phases that constitute a Rx ISI bit period.

The optimum sample position for an ideal ISI bit period is 2 isi_pclk cycles after the ISI bit boundary sample, which should result in a data sample close to the centre of the ISI bit period. rx_sample is asserted during the rx2 state to indicate a valid ISI data sample on rx_bit, unless the bit should be stripped when flagged by the bit stuffing statemachine, in which case rx_sample is not asserted during rx2 and the bit is not written to the FIFO. When edge is asserted, it resets the Rx cycle to the rx0 state, from any rx state. This is how the isi_sie tracks the phase of the incoming data. The Rx cycle will cycle through states rx0->rx4 until edge is asserted to reset the sample phase, or a tx_req is asserted indicating that the ISI needs to transmit.

Due to the 5 times oversampling a maximum phase error of 0.4 of an ISI bit period (2 isi_pclk cycles out of 5) can be tolerated.

NOTE: All statemachine signals are assumed to be ‘0’ unless otherwise stated.

An example of the Tx data generation mechanism is detailed in FIG. 46. tx_req and fifo_wr_tx are driven by the framer block.

An example of the Rx data sampling functional timing is detailed in FIG. 47. The dashed lines on the ser_data_in_ff signal indicate where the Rx/Tx statemachine perceived the bit boundary to be, based on the phase of the last ISI bit boundary. It can be seen that data is sampled during the same phase as the previous bit was, in the absence of a transition.

12.4.6.2.3 SIE Rx/Tx FIFO

The Rx/Tx FIFO is a 7×1 bit synchronous look-ahead FIFO that is shared for Tx and Rx operations. It is required to absorb any Rx/Tx latency caused by bit stripping/stuffing on a per ISI line basis, i.e. some ISI lines may require bit stripping/stuffing during an ISI bit period while the others may not, which would lead to a loss of synchronization between the data of the different ISI lines, if a FIFO were not present in each isi_sie.

The basic functional components of the FIFO are detailed in FIG. 48. tx_ready is driven by the Rx/Tx statemachine and selects which signals control the read and write operations. tx_ready=1 during ISI transmission and selects the fifo_*tx control and data signals. tx_ready=0 during ISI reception and selects the fifo_*rx control and data signals. fifo_reset is driven by the Rx/Tx statemachine. It is active high and resets the FIFO and associated logic before/after transmitting a packet to discard any residual data.

The size of the FIFO is based on the maximum bit stuffing frequency and the size of the shift register used to segment/re-assemble the multiple serial streams in the ISI framing logic. The maximum bit stuffing frequency is every 7 consecutive ones or zeroes. The shift register used is 32 bits wide. This implies that the maximum number of stuffed bits encountered in the time it takes to fill/empty the shift register if 4. This would suggest that 4×1 bit would be the minimum ideal size of the FIFO. However it is necessary to allow for different skew and phase error between the ISI lines, hence a 7×1 bit FIFO.

The FIFO is controlled by the isi_sie during packet reception and is controlled by the isi_frame block during packet transmission. This is illustrated in FIG. 49. The signal tx_ready selects which mode the FIFO control signals operate in. When tx_ready=0, i.e. Rx mode, the isi_sie control signals rx_sample, fifo_rd_rx and ser_data_in_ff are selected. When tx_ready=1, i.e. Tx mode, the sie_frame control signals fifo_wr_tx, fifo_rd_tx and fifo_wr_data_tx are selected.

12.4.6.3 Bit Stuffing

Programmable bit stuffing is implemented in the isi_sie. This is to allow the system to determine the amount of bit stuffing necessary for a specific ISI system devices. It is unlikely that bit stuffing would be required in a system using a 100 ppm rated crystal. However, a programmable bit stuffing implementation is much more versatile and robust.

The bit stuffing logic consists of a counter and a statemachine that track the number of consecutive ones or zeroes that are transmitted or received and flags the Rx/Tx statemachine when the bit stuffing limit has been reached. The counter, stuff_count, is a 7 bit counter, which decrements when rx_sample is asserted on a Rx cycle or when fifo_rd_tx is asserted on a Tx cycle. The upper 4 bits of stuff_count are loaded with isi_bit_stuff_rate. The lower 3 bits of stuff_count are always loaded with b111, i.e. for isi_bit_stuff_rate=b000, the counter would be loaded with b0000111. This is to prevent bit stuffing for less than 7 consecutive ones or zeroes. This allows the bit stuffing limit to be set in the range 7->127 consecutive ones or zeroes.

NOTE: It is extremely important that a change in the bit stuffing rate, isi_bit_stuff_rate, is carefully coordinated between ISI devices in a system. It is obvious that ISI devices will not be able to communicate reliably with each other with different bit stuffing settings. It is recommended that all ISI devices in a system default to the safest bit stuffing rate (isi_bit_stuff_rate=b000) at reset. The system can then co-ordinate the change to an optimum bit stuffing rate.

The ISI bit stuffing statemachine Tx cycle is shown in FIG. 50. The counter is loaded when stuff_count_load is asserted.

NOTE: All statemachine signals are assumed to be ‘0’ unless otherwise stated.

The ISI bit stuffing statemachine Rx cycle is shown in FIG. 51. It should be noted that the statemachine enters the strip state when stuff_count=0x2. This is because the statemachine can only transition to rx0 or rx1 when rx_sample is asserted as it needs to be synchronized to changes in sampling phase introduced by the Rx/Tx statemachine. Therefore a one or a zero has already been sampled by the time it enters rx0 or rx1. This is not the case for the Tx cycle, as it will always have a stable 5 isi_pclk cycles per bit period and relies purely on the data value when entering tx0 or tx1. The Tx cycle therefore enters stuff1 or stuff0 when stuff_count=0x1.

NOTE: All statemachine signals are assumed to be ‘0’ unless otherwise stated.

12.4.6.4 ISI Framing and CRC Sub-Block (isi_frame)

12.4.6.4.1 CRC Generation/Checking

A Cyclic Redundancy Checksum (CRC) is calculated over all fields except the start and stop fields for each long or ping packet transmitted. The receiving ISI device will perform the same calculation on the received packet to verify the integrity of the packet. The procedure used in the CRC generation/checking is the same as the Frame Checking Sequence (FCS) procedure used in HDLC, detailed in ITU-T Recommendation T30[39].

For generation/checking of the CRC field, the shift register illustrated in FIG. 52 is used to perform the modulo 2 division on the packet contents by the polynomial G(x)=x¹⁶+x¹²+x⁵+1.

To generate the CRC for a transmitted packet, where T(x)=[Packet Descriptor field, Address field, Data Payload field] (a ping packet will not contain a data payload field).

-   -   Set the shift register to 0xFFFF.     -   Shift T(x) through the shift register, LSB first. This can occur         in parallel with the packet transmission.     -   Once the each bit of T(x) has been shifted through the register,         it will contain the remainder of the modulo 2 division         T(x)/G(x).     -   Perform a ones complement of the register contents, giving the         CRC field which is transmitted MSB first, immediately following         the last bit of M(x         -   To check the CRC for a received packet, where R(x)=[Packet             Descriptor field, Address field, Data Payload field, CRC             field] (a ping packet will not contain a data payload             field).         -   Set the shift register to 0xFFFF.         -   Shift R(x) through the shift register, LSB first. This can             occur in parallel with the packet reception.         -   Once each bit of the packet has been shifted through the             register, it will contain the remainder of the modulo 2             division R(x)/G(x).         -   The remainder should equal b0001110100001111, for a packet             without errors.             12.5 CTRL (Control Sub-Block)             12.5.1 Overview

The CTRL is responsible for high level control of the SCB sub-blocks and coordinating access between them. All control and status registers for the SCB are contained within the CTRL and are accessed via the CPU interface. The other major components of the CTRL are the SCB Map logic and the DMA Manager logic.

12.5.2 SCB Mapping

In order to support maximum flexibility when moving data through a multi-SoPEC system it is possible to map any USB endpoint onto either DMAChannel within any SoPEC in the system. The SCB map, and indeed the SCB itself is based around the concept of an ISIId and an ISISubId. Each SoPEC in the system has a unique ISIId and two ISISubIds, namely ISISubId0 and ISISubId1. We use the convention that ISISubId0 corresponds to DMAChannel0 in each SoPEC and ISISubId1 corresponds to DMAChannel1. The naming convention for the ISIId is shown in Table 35 below and this would correspond to a multi-SoPEC system such as that shown in FIG. 27. We use the term ISIId instead of SoPECId to avoid confusion with the unique ChipID used to create the SoPEC_id and SoPEC_id_key (see chapter 17 and [9] for more details).

TABLE 35 ISIId naming convention ISIId SoPEC to which it refers 0-14 Standard device ISIIds (0 is the power-on reset value) 15 Broadcast ISIId

The combined ISIId and ISISubId therefore allows the ISI to address DMAChannel0 or DMAChannel1 on any SoPEC device in the system. The ISI, DMA manager and SCB map hardware use the ISIId and ISISubId to handle the different data streams that are active in a multi-SoPEC system as does the software running on the CPU of each SoPEC. In this document we will identify DMAChannels as ISIx.y where x is the ISIId and y is the ISISubId. Thus ISI2.1 refers to DMAChannel1 of ISISlave2. Any data sent to a broadcast channel, i.e. ISI15.0 or ISI15.1, are received by every ISI device in the system including the ISIMaster (which may be an ISI-Bridge). The USB device controller and software stacks however have no understanding of the ISIId and ISISubId but the Silverbrook printer driver software running on the external host does make use of the ISIId and ISISubId. USB is simply used as a data transport—the mapping of USB device endpoints onto ISIId and SubId is communicated from the external host Silverbrook code to the SoPEC Silverbrook code through USB control (or possibly bulk data) messages i.e. the mapping information is simply data payload as far as USB is concerned. The code running on SoPEC is responsible for parsing these messages and configuring the SCB accordingly.

The use of just two DMAChannels places some limitations on what can be achieved without software intervention. For every SoPEC in the system there are more potential sources of data than there are sinks. For example an ISISlave could receive both control and data messages from the ISIMaster SoPEC in addition to control and data from the external host, either specifically addressed to that particular ISISlave or over the broadcast ISI channel. However all ISISlaves only have two possible data sinks, i.e. DMAChannelz0 and DMAChannel1. Another example is the ISIMaster in a multi-SoPEC system which may receive control messages from each SoPEC in addition to control and data information from the external host (e.g. over USB). In this case all of the control messages are in contention for access to DMAChannel0. We resolve these potential conflicts by adopting the following conventions:

1) Control messages may be interleaved in a memory buffer: The memory buffer that the DMAChannel0 points to should be regarded as a central pool of control messages. Every control message must contain fields that identify the size of the message, the source and the destination of the control message. Control messages may therefore be multiplexed over a DMAChannel which allows several control message sources to address the same DMAChannel. Furthermore, if SoPEC-type control messages contain source and destination fields it is possible for the external host to send control messages to individual SoPECs over the ISI15.0 broadcast channel.

2) Data messages should not be interleaved in a memory buffer: As data messages are typically part of a much larger block of data that is being transferred it is not possible to control their contents in the same manner as is possible with the control messages. Furthermore we do not want the CPU to have to perform reassembly of data blocks. Data messages from different sources cannot be interleaved over the same DMAChannel—the SCB map must be reconfigured each time a different data source is given access to the DMAChannel.

3) Every reconfiguration of the SCB map requires the exchange of control messages: SoPEC's SCB map reset state is shown in Table and any subsequent modifications to this map require the exchange of control messages between the SoPEC and the external host. As the external host is expected to control the movement of data in any SoPEC system it is anticipated that all changes to the SCB map will be performed in response to a request from the external host. While the SoPEC could autonomously reconfigure the SCB map (this is entirely up to the software running on the SoPEC) it should not do so without informing the external host in order to avoid data being misrouted.

An example of the above conventions in operation is worked through in section 12.5.2.3.

12.5.2.1 SCB Map Rules

The operation of the SCB map is described by these 2 rules:

Rule 1: A packet is routed to the DMA manager if it originates from the USB device core and has an ISIId that matches the local SoPEC ISIId.

Rule 2: A packet is routed to the ISI if it originates from the CPU or has an ISIId that does not match the local SoPEC ISIId.

If the CPU erroneously addresses a packet to the ISIId contained in the ISIId register (i.e. the ISIId of the local SoPEC) then that packet will be transmitted on the ISI rather than be sent to the DMA manager. While this will usually cause an error on the ISI there is one situation where it could be beneficial, namely for initial dialog in a 2 SoPEC system as both devices come out of reset with an ISIId of 0.

12.5.2.2 External Host to ISIMaster SoPEC Communication

Although the SCB map configuration is independent of ISIMaster status, the following discussion on SCB map configurations assumes the ISIMaster is a SoPEC device rather than an ISI bridge chip, and that only a single USB connection to the external host is present. The information should apply broadly to an ISI-Bridge but we focus here on an ISIMaster SoPEC for clarity.

As the ISIMaster SoPEC represents the printer device on the PC USB bus it is required by the USB specification to have a dedicated control endpoint, EP0. At boot time the ISIMaster SoPEC will also require a bulk data endpoint to facilitate the transfer of program code from the external host. The simplest SCB map configuration, i.e. for a single stand-alone SoPEC, is sufficient for external host to ISIMaster SoPEC communication and is shown in Table 36.

TABLE 36 Single SoPEC SCB map configuration Source Sink EP0 ISI0.0 EP1 ISI0.1 EP2 nc EP3 nc EP4 nc

In this configuration all USB control information exchanged between the external host and SoPEC over EP0 (which is the only bidirectional USB endpoint). SoPEC specific control information (printer status, DNC info etc.) is also exchanged over EP0.

All packets sent to the external host from SoPEC over EP0 must be written into the DMA mapped EP buffer by the CPU (LEON-PC dataflow in FIG. 29). All packets sent from the external host to SoPEC are placed in DRAM by the DMA Manager, where they can be read by the CPU (PC-DIU dataflow in FIG. 29). This asymmetry is because in a multi-SoPEC environment the CPU will need to examine all incoming control messages (i.e. messages that have arrived over DMAChannel0) to ascertain their source and destination (i.e. they could be from an ISISlave and destined for the external host) and so the additional overhead in having the CPU move the short control messages to the EP0 FIFO is relatively small. Furthermore we wish to avoid making the SCB more complicated than necessary, particularly when there is no significant performance gain to be had as the control traffic will be relatively low bandwidth.

The above mechanisms are appropriate for the types of communication outlined in sections 12.1.2.1.1 through 12.1.2.1.4

12.5.2.3 Broadcast Communication

The SCB configuration for broadcast communication is also the default, post power-on reset, configuration for SoPEC and is shown in Table 37.

TABLE 37 Default SoPEC SCB map configuration Source Sink EP0 ISI0.0 EP1 ISI0.1 EP2 ISI15.0 EP3 ISI15.1 EP4 ISI1.1

USB endpoints EP2 and EP3 are mapped onto ISISubID0 and ISISubId1 of ISIId15 (the broadcast ISIId channel). EP0 is used for control messages as before and EP1 is a bulk data endpoint for the ISIMaster SoPEC. Depending on what is convenient for the boot loader software, EP1 may or may not be used during the initial program download, but EP1 is highly likely to be used for compressed page or other program downloads later. For this reason it is part of the default configuration. In this setup the USB device configuration will take place, as it always must, by exchanging messages over the control channel (EP0).

One possible boot mechanism is where the external host sends the bootloader1 program code to all SoPECs by broadcasting it over EP3. Each SoPEC in the system then authenticates and executes the bootloader1 program. The ISIMaster SoPEC then polls each ISISlave (over the ISIx.0 channel). Each ISISlave ascertains its ISIId by sampling the particular GPIO pins required by the bootloader1 and reporting its presence and status back to the ISIMaster. The ISIMaster then passes this information back to the external host over EP0. Thus both the external host and the ISIMaster have knowledge of the number of SoPECs, and their ISIIds, in the system. The external host may then reconfigure the SCB map to better optimise the SCB resources for the particular multi-SoPEC system. This could involve simplifying the default configuration to a single SoPEC system or remapping the broadcast channels onto DMAChannels in individual ISISlaves.

The following steps are required to reconfigure the SCB map from the configuration depicted in Table to one where EP3 is mapped onto ISI1.0:

-   1) The external host sends a control message(s) to the ISIMaster     SoPEC requesting that USB EP3 be remapped to ISI1.0 -   2) The ISIMaster SoPEC sends a control message to the external host     informing it that EP3 has now been mapped to ISI1.0 (and therefore     the external host knows that the previous mapping of ISI15.1 is no     longer available through EP3). -   3) The external host may now send control messages directly to     ISISlave1 without requiring any CPU intervention on the ISIMaster     SoPEC     12.5.2.4 External Host to ISISlave SoPEC Communication

If the ISIMaster is configured correctly (e.g. when the ISIMaster is a SoPEC, and that SoPEC's SCB map is configured correctly) then data sent from the external host destined for an ISISlave will be transmitted on the ISI with the correct address. The ISI automatically forwards any data addressed to it (including broadcast data) to the DMA channel with the appropriate ISISubId. If the ISISlave has data to send to the external host it must do so by sending a control message to the ISIMaster identifying the external host as the intended recipient. It is then the ISIMaster's responsibility to forward this message to the external host.

With this configuration the external host can communicate with the ISISlave via broadcast messages only and this is the mechanism by which the bootloader1 program is downloaded. The ISISlave is unable to communicate with the external host (or the ISIMaster) until the bootloader1 program has successfully executed and the ISISlave has determined what its ISIId is. After the bootloader1 program (and possibly other programs) has executed the SCB map of the ISIMaster may be reconfigured to reflect the most appropriate topology for the particular multi-SoPEC system it is part of.

All communication from an ISISlave to external host is either achieved directly (if there is a direct USB connection present for example) or by sending messages via the ISIMaster. The ISISlave can never initiate communication to the external host. If an ISISlave wishes to send a message to the external host via the ISIMaster it must wait until it is pinged by the ISIMaster and then send a the message in a long packet addressed to the ISIMaster. When the ISIMaster receives the message from the ISISlave it first examines it to determine the intended destination and will then copy it into the EP0 FIFO for transmission to the external host. The software running on the ISIMaster is responsible for any arbitration between messages from different sources (including itself) that are all destined for the external host.

The above mechanisms are appropriate for the types of communication outlined in sections 12.1.2.1.5 and 12.1.2.1.6.

12.5.2.5 ISIMaster to ISISlave Communication

All ISIMaster to ISISlave communication takes place over the ISI. Immediately after reset this can only be by means of broadcast messages. Once the bootloader1 program has successfully executed on all SoPECs in a multi-SoPEC system the ISIMaster can communicate with each SoPEC on an individual basis.

If an ISISlave wishes to send a message to the ISIMaster it may do so in response to a ping packet from the ISIMaster. When the ISIMaster receives the message from the ISISlave it must interpret the message to determine if the message contains information required to be sent to the external host. In the case of the ISIMaster being a SoPEC, software will transfer the appropriate information into the EP0 FIFO for transmission to the external host.

The above mechanisms are appropriate for the types of communication outlined in sections 12.1.2.3.3 and 12.1.2.3.4.

12.5.2.6 ISISlave to ISISlave Communication

ISISlave to ISISlave communication is expected to be limited to two special cases: (a) when the PrintMaster is not the ISIMaster and (b) when a storage SoPEC is used. When the PrintMaster is not the ISIMaster then it will need to send control messages (and receive responses to these messages) to other ISISlaves. When a storage SoPEC is present it may need to send data to each SoPEC in the system. All ISISlave to ISISlave communication will take place in response to ping messages from the ISIMaster.

12.5.2.7 Use of the SCB Map in an ISISlave with a External Host Connection

After reset any SoPEC (regardless of ISIMaster/Slave status) with an active USB connection will route packets from EP0,1 to DMA channels 0,1 because the default SCB map is to map EP0 to ISIId0.0 and EP1 to ISIId0.1 and the default ISIId is 0. At some later time the SoPEC learns its true ISIId for the system it is in and re-configures its ISIId and SCB map registers accordingly. Thus if the true ISIId is 3 the external host could reconfigure the SCB map so that EP0 and EP1 (or any other endpoints for that matter) map to ISIId3.0 and 3.1 respectively. The co-ordination of the updating of the ISIId registers and the SCB map is a matter for software to take care of. While the AutoMasterEnable bit of the ISICntrl register is set the external host must not send packets down EP2-4 of the USB connection to the device intended to be an ISISlave. When AutoMasterEnable has been cleared the external host may send data down any endpoint of the USB connection to the ISISlave.

The SCB map of an ISISlave can be configured to route packets from any EP to any ISIId.ISISubId oust as an ISIMaster can). As with an ISIMaster these packets will end up in the SCBTxBuffer but while an ISIMaster would just transmit them when it got a local access slot (from ping arbitration) the ISISlave can only transmit them in response to a ping. All this would happen without CPU intervention on the ISISlave (or ISIMaster) and as long as the ping frequency is sufficiently high it would enable maximum use of the bandwidth on both USB buses.

12.5.3 DMA Manager

The DMA manager manages the flow of data between the SCB and the embedded DRAM. Whilst the CPU could be used for the movement of data in SoPEC, a DMA manager is a more efficient solution as it will handle data in a more predictable fashion with less latency and requiring less buffering. Furthermore a DMA manager is required to support the ISI transfer speed and to ensure that the SoPEC could be used with a high speed ISI-Bridge chip in the future.

The DMA manager utilizes 2 write channels (DMAChannel0, DMAChannel1) and 1 read/write channel (DMAChannel2) to provide 2 independent modes of access to DRAM via the DIU interface:

-   -   USBD/ISI type access.     -   USBH type access.

DIU read and write access is in bursts of 4×64 bit words. Byte aligned write enables are provided for write access. Data for DIU write accesses will be read directly from the buffers contained in the respective SCB sub-blocks. There is no internal SCB DMA buffer. The DMA manager handles all issues relating to byte/word/longword address alignment, data endianness and transaction scheduling. If a DMA channel is disabled during a DMA access, the access will be completed. Arbitration will be performed between the following DIU access requests:

-   -   USBD write request.     -   ISI write request.     -   USBH write request.     -   USBH read request.

DMAChannel0 will have absolute priority over any DMA requesters. In the absence of DMAChannel0 DMA requests, arbitration will be performed in a round robin manner, on a per cycle basis over the other channels.

12.5.3.1 DMA Effective Bandwidth

The DIU bandwidth available to the DMA manager must be set to ensure adequate bandwidth for all data sources, to avoid back pressure on the USB and the ISI. This is achieved by setting the output (i.e. DIU) bandwidth to be greater than the combined input bandwidths (i.e. USBD+USBH+ISI). The required bandwidth is expected to be 160 Mbits/s (1 bit/cycle @ 160 MHz). The guaranteed DIU bandwidth for the SCB is programmable and may need further analysis once there is better knowledge of the data throughput from the USB IP cores.

12.5.3.2 USBDIISI DMA Access

The DMA manager uses the two independent unidirectional write channels for this type of DMA access, one for each ISISubID, to control the movement of data. Both DMAChannel0 and DMAChannel1 only support write operation and can transfer data from any USB device DMA mapped EP buffer and from the ISI receive buffer to separate circular buffers in DRAM, corresponding to each DMA channel.

While the DMA manager performs the work of moving data the CPU controls the destination and relative timing of data flows to and from the DRAM. The management of the DRAM data buffers requires the CPU to have accurate and timely visibility of both the DMA and PEP memory usage. In other words when the PEP has completed processing of a page band the CPU needs to be aware of the fact that an area of memory has been freed up to receive incoming data. The management of these buffers may also be performed by the external host.

12.5.3.2.1 Circular Buffer Operation

The DMA manager supports the use of circular buffers for both DMAChannels. Each circular buffer is controlled by 5 registers: DMAnBottomAdr, DMAnTopAdr, DMAnMaxAdr, DMAnCurrWPtr and DMAnIntAdr. The operation of the circular buffers is shown in FIG. 53 below.

Here we see two snapshots of the status of a circular buffer with (b) occurring sometime after (a) and some CPU writes to the registers occurring in between (a) and (b). These CPU writes are most likely to be as a result of a finished band interrupt (which frees up buffer space) but could also have occurred in a DMA interrupt service routine resulting from DMAnIntAdr being hit. The DMA manager will continue filling the free buffer space depicted in (a), advancing the DMAnCurrWPtr after each write to the DIU. Note that the DMACurrWPtr register always points to the next address the DMA manager will write to. When the DMA manager reaches the address in DMAnIntAdr (i.e. DMACurrWPtr=DMAnIntAdr) it will generate an interrupt if the DMAnIntAdrMask bit in the DMAMask register is set. The purpose of the DMAnintAdr register is to alert the CPU that data (such as a control message or a page or band header) has arrived that it needs to process. The interrupt routine servicing the DMA interrupt will change the DMAnintAdr value to the next location that data of interest to the CPU will have arrived by.

In the scenario shown in FIG. 53 the CPU has determined (most likely as a result of a finished band interrupt) that the filled buffer space in (a) has been freed up and is therefore available to receive more data. The CPU therefore moves the DMAnMaxAdr to the end of the section that has been freed up and moves the DMAnIntAdr address to an appropriate offset from the DMAnMaxAdr address. The DMA manager continues to fill the free buffer space and when it reaches the address in DMAnTopAdr it wraps around to the address in DMAnBottomAdr and continues from there. DMA transfers will continue indefinitely in this fashion until the DMA manager reaches the address in the DMAnMaxAdr register.

The circular buffer is initialized by writing the top and bottom addresses to the DMAnTopAdr and DMAnBottomAdr registers, writing the start address (which does not have to be the same as the DMAnBottomAdr even though it usually will be) to the DMAnCurrWPtr register and appropriate addresses to the DMAnIntAdr and DMAnMaxAdr registers. The DMA operation will not commence until a 1 has been written to the relevant bit of the DMAChanEn register.

While it is possible to modify the DMAnTopAdr and DMAnBottomAdr registers after the DMA has started it should be done with caution. The DMAnCurrWPtr register should not be written to while the DMAChannel is in operation. DMA operation may be stalled at any time by clearing the appropriate bit of the DMAChanEn register or by disabling an SCB mapping or ISI receive operation.

12.5.3.2.2 Non-Standard Buffer Operation

The DMA manager was designed primarily for use with a circular buffer. However because the DMA pointers are tested for equality (i.e. interrupts generated when DMAnCurrWPtr=DMAIntAdr or DMAnCurrWPtr=DMAMaxAdr) and no bounds checking is performed on their values (i e. neither DMAnIntAdr nor DMAnMaxAdr are checked to see if they lie between DMAnBottomAdr and DMAnTopAdr) a number of non-standard buffer arrangements are possible. These include:

-   -   Dustbin buffer: If DMAnBottomAdr, DMAnTopAdr and DMAnCurrWPtr         all point to the same location and both DMAnIntAdr and         DMAnMaxAdr point to anywhere else then all data for that DMA         channel will be dumped into the same location without ever         generating an interrupt. This is the equivalent to writing to         /dev/null on Unix systems.     -   Linear buffer: If DMAnMaxAdr and DMAnTopAdr have the same value         then the DMA manager will simply fill from DMAnBottomAdr to         DMAnTopAdr and then stop. DMAnIntAdr should be outside this         buffer or have its interrupt disabled.         12.5.3.3 USBH DMA Access

The USBH requires DMA access to DRAM in to provide a communication channel between the USB HC and the USB HCD via a shared memory resource. The DMA manager uses two independent channels for this type of DMA access, one for reads and one for writes. The DRAM addresses provided to the DIU interface are generated based on addresses defined in the USB HC core operational registers, in USBH section 12.3.

12.5.3.4 Cache Coherency

As the CPU will be processing some of the data transferred (particularly control messages and page/band headers) into DRAM by the DMA manager, care needs to be taken to ensure that the data it uses is the most recently transferred data. Because the DMA manager will be updating the circular buffers in DRAM without the knowledge of the cache controller logic in the LEON CPU core the contents of the cache can become outdated. This situation can be easily handled by software, for example by flushing the relevant cache lines, and so there is no hardware support to enforce cache coherency.

12.5.4 ISI Transmit Buffer Arbitration

The SCB control logic will arbitrate access to the ISI transmit buffer (ISITxBuffer) interface on the ISI block. There are two sources of ISI Tx packets:

-   -   CPUISITxBuffer, contained in the SCB control block.     -   ISI mapped USB EP OUT buffers, contained in the USB device         block.

This arbitration is controlled by the ISITxBuffArb register which contains a high priority bit for both the CPU and the USB. If only one of these bits is set then the corresponding source always has priority. Note that if the CPU is given absolute priority over the USB, then the software filling the ISI transmit buffer needs to ensure that sufficient USB traffic is allowed through. If both bits of the ISITxBufferArb have the same value then arbitration will take place on a round robin basis. The control logic will use the USBEPnDest registers, as it will use the CPUISITxBuffCntrl register, to determine the destination of the packets in these buffers. When the ISITxBuffer has space for a packet, the SCB control logic will immediately seek to refill it. Data will be transferred directly from the CPUISITxBuffer and the ISI mapped USB EP OUT buffers to the ISITxBuffer without any intermediate buffering.

As the speed at which the ISITxBuffer can be emptied is at least 5 times greater than it can be filled by USB traffic, the ISI mapped USB EP OUT buffers should not overflow using the above scheme in normal operation. There are a number of scenarIOs which could lead to the USB EPs being temporarily blocked such as the CPU having priority, retransmissions on the ISI bus, channels being enabled (ChannelEn bit of the USBEPnDest register) with data already in their associated endpoint buffers or short packets being sent on the USB. Care should be taken to ensure that the USB bandwidth is efficiently utilised at all times.

12.5.5 Implementation

12.5.5.1 CTRL Sub-Block Partition

-   -   Block Diagram     -   Definition of I/Os         12.5.5.2 SCB Configuration Registers

The SCB register map is listed in Table 38. Registers are grouped according to which SCB sub-block their functionality is associated. All configuration registers reside in the CTRL sub-block. The Reset values in the table indicates the 32 bit hex value that will be returned when the CPU reads the associated address location after reset. All Registers pre-fixed with Hc refer to Host Controller Operational Registers, as defined in the OHCI Spec[19].

The SCB will only allow supervisor mode accesses to data space (i.e. cpu_acode[1:0]=b11). All other accesses will result in scb_cpu_berr being asserted.

TDB: Is read access necessary for ISI Rx/Tx buffers? Could implement the ISI interface as simple FIFOs as opposed to a memory interface.

TABLE 38 SCB control block configuration registers Address Offset from SCB_base Register #Bits Reset Description CTRL 0x000 SCBResetN 4 0x0000000F SCB software reset. Allows individual sub-blocks to be reset separately or together. Once a reset for a block has been initiated, by writing a 0 to the relevant register field, it can not be suppressed. Each field will be set after reset. Writing 0x0 to the SCBReset register will have the same effect as CPR generated hardware reset. 0x004 SCBGo 2 0x00000000 SCB Go. Allows the ISI and CTRL sub-blocks to be selected separately or together. When go is de-asserted for a particular sub-block, its statemachines are reset to their idle states and its interface signals are de-asserted. The sub-block counters and configuration registers retain their values. When go is asserted for a particular sub-block, its counters are reset. The sub-block configuration registers retain their values, i.e. they don't get reset. The sub-block statemachines and interface signals will return to their normal mode of operation. The CTRL field should be de-asserted before disabling the clock from any part of the SCB to avoid erroneous SCB DMA requests when the clock is enabled again. NOTE: This functionality has not been provided for the USBH and USBD sub- blocks because of the USB IP cores that they contain. We do not have direct control over the IP core statemachines and counters, and it would cause unpredictable behaviour if the cores were disabled in this way during operation. 0x008 SCBWakeupEn 2 0x00000000 USB/ISI WakeUpEnable register 0x00C SCBISITxBufferArb 2 0x00000000 ISI transmit buffer access priority register. 0x010 SCBDebugSel[11:2] 10 0x00000000 SCB Debug select register. 0x014 USBEP0Dest 7 0x00000020 This register determines which of the data sinks the data arriving in EP0 should be routed to. 0x018 USBEP1Dest 7 0x00000021 Data sink mapping for USB EP1 0x01C USBEP2Dest 7 0x0000003E Data sink mapping for USB EP2 0x020 USBEP3Dest 7 0x0000003F Data sink mapping for USB EP3 0x024 USBEP4Dest 7 0x00000023 Data sink mapping for USB EP4 0x028 DMA0BottomAdr[21:5] 17 DMAChannel0 bottom address register. 0x02C DMA0TopAdr[21:5] 17 DMAChannel0 top address register. 0x030 DMA0CurrWPtr[21:5] 17 DMAChannel0 current write pointer. 0x034 DMA0IntAdr[21:5] 17 DMAChannel0 interrupt address register. 0x038 DMA0MaxAdr[21:5] 17 DMAChannel0 max address register. 0x03C DMA1BottomAdr[21:5] 17 As per DMA0BottomAdr. 0x040 DMA1TopAdr[21:5] 17 As per DMA0TopAdr. 0x044 DMA1CurrWPtr[21:5] 17 As per DMA0CurrWPtr. 0x048 DMA1IntAdr[21:5] 17 As per DMA0IntAdr. 0x04C DMA1MaxAdr[21:5] 17 As per DMA0MaxAdr. 0x050 DMAAccessEn 3 0x00000003 DMA access enable. 0x054 DMAStatus 4 0x00000000 DMA status register. 0x058 DMAMask 4 0x00000000 DMA mask register. 0x05C-0x098 CPUISITxBuff[7:0] 32 × 8 n/a CPU ISI transmit buffer. 32-byte packet buffer, containing the payload of a CPU sourced packet destined for transmission over the ISI. The CPU has full write access to the CPUISITxBuff. NOTE: The CPU does not have read access to CPUISITxBuff. This is because the CPU is the source of the data and to avoid arbitrating read access between the CPU and the CTRL sub-block. Any CPU reads from this address space will return 0x00000000. 0x09C CPUISITxBuffCtrl 9 0x00000000 CPU ISI transmit buffer control register. USBD 0x100 USBDIntStatus 19 0x00000000 USBD Interrupt event status register. 0x104 USBDISIFIFOStatus 16 0x00000000 USBD ISI mapped OUT EP packet FIFO status register. 0x108 USBDDMA0FIFOStatus 8 0x00000000 USBD DMAChannel0 mapped OUT EP packet FIFO status register. 0x10C USBDDMA1FIFOStatus 8 0x00000000 USBD DMAChannel1 mapped OUT EP packet FIFO status register. 0x110 USBDResume 1 0x00000000 USBD core resume register. 0x114 USBDSetup 4 0x00000000 USBD setup/configuration register. 0x118-0x154 USBDEp0InBuff[15:0]  32 × 16 n/a USBD EP0-IN buffer. 64-byte packet buffer in the, containing the payload of a USB packet destined for EP0-IN. The CPU has full write access to the USBDEp0InBuff. NOTE: The CPU does not have read access to USBDEp0InBuff. This is because the CPU is the source of the data and to avoid arbitrating read access between the CPU and the USB device core. Any CPU reads from this address space will return 0x00000000. 0x158 USBDEp0InBuffCtrl 1 0x00000000 USBD EP0-IN buffer control register. 0x15C-0x198 USBDEp5InBuff[15:0]  32 × 16 n/a USBD EP5-IN buffer. As per USBDEp0InBuff. 0x19C USBDEp5InBuffCtrl 1 0x00000000 USBD EP5-IN buffer control register. 0x1A0 USBDMask 19 0x00000000 USBD interrupt mask register. 0x1A4 USBDDebug 30 0x00000000 USBD debug register. USBH 0x200 HcRevision Refer to [19] for #Bits, Reset, Description. 0x204 HcControl Refer to [19] for #Bits, Reset, Description. 0x208 HcCommandStatus Refer to [19] for #Bits, Reset, Description. 0x20C HcInterruptStatus Refer to [19] for #Bits, Reset, Description. 0x210 HcInterruptEnable Refer to [19] for #Bits, Reset, Description. 0x214 HcInterruptDisable Refer to [19] for #Bits, Reset, Description. 0x218 HcHCCA Refer to [19] for #Bits, Reset, Description. 0x21C HcPeriodCurrentED Refer to [19] for #Bits, Reset, Description. 0x220 HcControlHeadED Refer to [19] for #Bits, Reset, Description. 0x224 HcControlCurrentED Refer to [19] for #Bits, Reset, Description. 0x228 HcBulkHeadED Refer to [19] for #Bits, Reset, Description. 0x22C HcBulkCurrentED Refer to [19] for #Bits, Reset, Description. 0x230 HcDoneHead Refer to [19] for #Bits, Reset, Description. 0x234 HcFmInterval Refer to [19] for #Bits, Reset, Description. 0x238 HcFmRemaining Refer to [19] for #Bits, Reset, Description. 0x23C HcFmNumber Refer to [19] for #Bits, Reset, Description. 0x240 HcPeriodicStart Refer to [19] for #Bits, Reset, Description. 0x244 HcLSTheshold Refer to [19] for #Bits, Reset, Description. 0x248 HcRhDescriptorA Refer to [19] for #Bits, Reset, Description. 0x24C HcRhDescriptorB Refer to [19] for #Bits, Reset, Description. 0x250 HcRhStatus Refer to [19] for #Bits, Reset, Description. 0x254 HcRhPortStatus[1] Refer to [19] for #Bits, Reset, Description. 0x258 USBHStatus 3 0x00000000 USBH status register. 0x25C USBHMask 2 0x00000000 USBH interrupt mask register. 0x260 USBHDebug 2 0x00000000 USBH debug register. ISI 0x300 ISICntrl 4 0x0000000B ISI Control register 0x304 ISIId 4 0x00000000 ISIId for this SoPEC. 0x308 ISINumRetries 4 0x00000002 Number of ISI retransmissions register. 0x30C ISIPingSchedule0 15 0x00000000 ISI Ping schedule 0 register. 0x310 ISIPingSchedule1 15 0x00000000 ISI Ping schedule 1 register. 0x314 ISIPingSchedule2 15 0x00000000 ISI Ping schedule 2 register. 0x318 ISITotalPeriod 4 0x0000000F Reload value of the ISITotalPeriod counter. 0x31C ISILocalPeriod 4 0x0000000F Reload value of the ISILocalPeriod counter. 0x320 ISIIntStatus 4 0x00000000 ISI interrupt status register. 0x324 ISITxBuffStatus 27 0x00000000 ISI Tx buffer status register. 0x328 ISIRxBuffStatus 27 0x00000000 ISI Rx buffer status register. 0x32C ISIMask 4 0x00000000 ISI Interrupt mask register. 0x330-0x34C ISITxBuffEntry0[7:0] 32 × 8 n/a ISI transmit Buff, packet entry #0. 32-byte packet entry in the ISITxBuff, containing the payload of an ISI Tx packet. CPU read access to ISITxBuffEntry0 is provided for observability only i.e. CPU reads of the ISITxBuffEntry0 do not alter the state of the buffer. The CPU does not have write access to the ISITxBuffEntry0. 0x350-0x36C ISITxBuffEntry1[7:0] 32 × 8 n/a ISI transmit Buff, packet entry #1. As per ISITxBuffEntry0. 0x370-0x38C ISIRxBuffEntry0[7:0] 32 × 8 n/a ISI receive Buff, packet entry #0. 32-byte packet entry in the ISIRxBuff, containing the payload of an ISI Rx packet. Note that the only error-free long packets are placed in the ISIRxBuffEntry0. Both ping and ACKs are consumed in the ISI. CPU access to ISIRxBuffEntry0 is provided for observability only i.e. CPU reads of the ISIRxBuffEntry0 do not alter the state of the buffer. 0x390-0x3AC ISIRxBuffEntry1[7:0] 32 × 8 n/a ISI receive Buff, packet entry #1. As per ISIRxBuffEntry0. 0x3B0 ISISubId0Seq 1 0x00000000 ISI sub ID 0 sequence bit register. 0x3B4 ISISubId1Seq 1 0x00000000 ISI sub ID 1 sequence bit register. 0x3B8 ISISubIdSeqMask 2 0x00000000 ISI sub ID sequence bit mask register. 0x3BC ISINumPins 1 0x00000000 ISI number of pins register. 0x3C0 ISITurnAround 4 0x0000000F ISI bus turn around register. 0x3C4 ISITShortReplyWin 5 0x0000001F ISI short packet reply window. 0x3C8 ISITLongReplyWin 9 0x000001FF ISI long packet reply window. 0x3CC ISIDebug 4 0x00000000 ISI debug register.

A detailed description of each register format follows. The CPU has full read access to all registers. Write access to the fields of each register is defined as:

-   -   Full: The CPU has full write access to the field, i.e. the CPU         can write a 1 or a 0 to each bit.     -   Clear: The CPU can clear the field by writing a 1 to each bit.         Writing a 0 to this type of field will have no effect.     -   None: The CPU has no write access to the field, i.e. a CPU write         will have no effect on the field.         12.5.5.2.1 SCBResetN

TABLE 39 SCBResetN register format Field Name Bit(s) write access Description CTRL 0 Full scb_ctrl sub-block reset. Setting this field will reset the SCB control sub-block logic, including all configuration registers. 0 = reset 1 = default state ISI 1 Full scb_isi sub-block reset. Setting this field will reset the ISI sub-block logic. 0 = reset 1 = default state USBH 2 Full scb_usbh sub-block reset. Setting this field will reset the USB host controller core and associated logic. 0 = reset 1 = default state USBD 3 Full scb_usbd sub-block reset. Setting this field will reset the USB device controller core and associated logic. 0 = reset 1 = default state 12.5.5.2.2 SCBGo

TABLE 40 SCBGo register format Field Name Bit(s) write access Description CTRL 0 Full scb_ctrl sub-block go. 0 = halted 1 = running ISI 1 Full scb_isi sub-block go. 0 = halted 1 = running 12.5.5.2.3 SCBWakeUpEn

This register is used to gate the propagation of the USB and ISI reset signals to the CPR block.

TABLE 41 SCBWakeUpEn register format Field Name Bit(s) write access Description USBWakeUpEn 0 Full usb_cpr_reset_n propagation enable. 1 = enable 0 = disable ISIWakeUpEn 1 Full isi_cpr_reset_n propagation enable. 1 = enable 0 = disable 12.5.5.2.4 SCBISITxBufferArb

This register determines which source has priority at the ISITxBuffer interface on the ISI block. When a bit is set priority is given to the relevant source. When both bits have the same value, arbitration will be performed in a round-robin manner.

TABLE 42 SCBISITxBufferArb register format write Field Name Bit(s) access Description CPUPriority 0 Full CPU priority 1 = high priority 0 = low priority USBPriority 1 Full USB priority 1 = high priority 0 = low priority 12.5.5.2.5 SCBDebugSel

Contains address of the register selected for debug observation as it would appear on cpu_adr. The contents of the selected register are output in the scb_cpu_data bus while cpu_scb_sel is low and scb_cpu_debug_valid is asserted to indicate the debug data is valid. It is expected that a number of pseudo-registers will be made available for debug observation and these will be outlined with the implementation details.

TABLE 43 SCBDebugSel register format write Field Name Bit(s) access Description CPUAdr 11:2 Full cpu_adr register address. 12.5.5.2.6 USBEPnDest

This register description applies to USBEP0Dest, USBEP1Dest, USBEP2Dest, USBEP3Dest, USBEP4Dest. The SCB has two routing options for each packet received, based on the DestISIId associated with the packets source EP:

-   -   To the DMA Manager     -   To the ISI

The SCB map therefore does not need special fields to identify the DMAChannels on the ISIMaster SoPEC as this is taken care of by the SCB hardware. Thus the USBEP0Dest and USBEP1Dest registers should be programmed with 0x20 and 0x21 (for ISI0.0 and ISI0.1) respectively to ensure data arriving on these endpoints is moved directly to DRAM.

TABLE 44 USBEPnDest register format Write Field Name Bit(s) access Description SequenceBit 0 Full Sequence bit for packets going from USBEPn to DestISIId.DestISISubId. Every CPU write to this register initialises the value of the sequence bit and this is subsequently updated by the ISI after every successful long packet transmission. DestISIId 4:1 Full Destination ISI ID. Denotes the ISIId of the target SoPEC as per Table DestISISubId 5 Full Destination ISI sub ID. Indicates which DMAChannel of the target SoPEC the endpoint is mapped onto: 0 = DMAChannel0 1 = DMAChannel1 ChannelEn 6 Full Communication channel enable bit for EPn. This enables/disables the communication channel for EPn. When disabled, the SCB will not accept USB packets addressed to EPn. 0 = Channel disabled 1 = Channel enabled

If the local SoPEC is connected to an external USB host, it is recommended that the EP0 communication channel should always remain enabled and mapped to DMAChannel0 on the local SoPEC, as this is intended as the primary control communication channel between the external USB host and the local SoPEC.

A SoPEC ISIMaster should map as many USB endpoints, under the control of the external host, as are required for the multi-SoPEC system it is part of. As already mentioned this mapping may be dynamically reconfigured.

12.5.5.2.7 DMAnBottomAdr

This register description applies to DMA0BottomAdr and DMA1BottomAdr.

TABLE 45 DMAnBottomAdr register format Field Name Bit(s) Write access Description DMAnBottomAdr 21:5 Full The 256-bit aligned DRAM address of the bottom of the circular buffer (inclusive) serviced by DMAChanneln 12.5.5.2.8 DMAnTopAdr

This register description applies to DMA0TopAdr and DMA 1 TopAdr.

TABLE 46 DMAnTopAdr register format Field Name Bit(s) Write access Description DMAnTopAdr 21:5 Full The 256-bit aligned DRAM address of the top of the circular buffer (inclusive) serviced by DMAChanneln 12.5.5.2.9 DMAnCurrWPtr

This register description applies to DMA0CurrWPtr and DMA1CurrWPtr.

TABLE 47 DMAnCurrWptr register format Write Field Name Bit(s) access Description DMAnCurrWPtr 21:5 Full The 256-bit aligned DRAM address of the next location DMAChannel0 will write to. This register is set by the CPU at the start of a DMA operation and dynamically updated by the DMA manager during the operation. 12.5.5.2.10 DMAIntAdr

This register description applies to DMA0IntAdr and DMA1IntAdr.

TABLE 48 DMAnIntAdr register format Write Bit(s) access Description DMAnIntAdr 21:5 Full The 256-bit aligned DRAM address of the location that will trigger an interrupt when reached by DMAChanneln buffer. 12.5.5.2.11 DMAnMaxAdr

This register description applies to DMA0MaxAdr and DMA1MaxAdr.

TABLE 49 DMAnMaxAdr register format Write Field Name Bit(s) access Description DMAnMaxAdr 21:5 Full The 256-bit aligned DRAM address of the last free location that in the DMAChanneln circular buffer. DMAChannel0 transfers will stop when it reaches this address. 12.5.5.2.12 DMAAccessEn

This register enables DMA access for the various requesters, on a per channel basis.

TABLE 50 DMAAccessEn register format Write Field Name Bit(s) access Description DMAChannel0En 0 Full DMA Channel #0 access enable. This uni-directional write channel is used by the USBD and the ISI. 1 = enable 0 = disable DMAChannel1En 1 Full As per USBDISI0En. DMAChannel2En 2 Full DMA Channel #2 access enable. This bi-directional read/write channel is used by the USBH. 1 = enable 0 = disable 12.5.5.2.13 DMAStatus

The status bits are not sticky bits i.e. they reflect the ‘live’ status of the channel. DMAChannelNntAdrHit and DMAChannelNMaxAdrHit status bits may only be cleared by writing to the relevant DMAnintAdr or DMAnMaxAdr register.

TABLE 51 DMAStatus register format Write Field Name Bit(s) access Description DMAChannel0IntAdrHit 0 None DMA channel #0 interrupt address hit. 1 = DMAChannel0 has reached the address contained in the DMA0IntAdr register. 0 = default state DMAChannel0MaxAdrHit 1 None DMA channel #0 max address hit. 1 = DMAChannel0 has reached the address contained in the DMA0MaxAdr register. 0 = default state DMAChannel1IntAdrHit 3 None As per DMAChannel0IntAdrHit. DMAChannel1MaxAdrHit 4 None As per DMAChannel0MaxAdrHit. 12.5.5.2.14 DMAMask Register

All bits of the DMAMask are both readable and writable by the CPU. The DMA manager cannot alter the value of this register. All interrupts are generated in an edge sensitive manner i.e. the DMA manager will generate a dma_icu_irq pulse each time a status bit goes high and its corresponding mask bit is enabled.

TABLE 52 DMAMask register format Write Field Name Bit(s) access Description DMAChannel0IntAdrHitIntEn 0 Full DMAChannel0IntAdrHit status interrupt enable. 1 = enable 0 = disable DMAChannel0MaxAdrHitIntEn 1 Full DMAChannel0MaxAdrHit status interrupt enable. 1 = enable 0 = disable DMAChannel1IntAdrHitIntEn 2 Full As per DMAChannel0IntAdrHitIntEn DMAChannel1MaxAdrHitIntEn 3 Full As per DMAChannel0MaxAdrHitIntEn 12.5.5.2.15 CPUISITxBuffCtrl Register

TABLE 53 CPUISITxBuffCtrl register format Write Field Name Bit(s) access Description PktValid 0 full This field should be set by the CPU to indicate the validity of the CPUISITxBuff contents. This field will be cleared by the SCB once the contents of the CPUISITxBuff has been copied to the ISITxBuff. NOTE: The CPU should not clear this field under normal operation. If the CPU clears this field during a packet transfer to the ISITxBuff, the transfer will be aborted - this is not recommended. 1 = valid packet. 0 = default state. PktDesc 3:1 full PktDesc field, as per Table , of the packet contained in the CPUISITxBuff. The CPU is responsible for maintaining the correct sequence bit value for each ISIId.ISISubId channel it communicates with. Only valid when CPU- ISITxBuffCtrl.PktValid = 1. DestISIId 7:4 full Denotes the ISIId of the target SoPEC as per Table . DestISISubId 8 full Indicates which DMAChannel of the target SoPEC the packet in the CPUISITxBuff is destined for. 1 = DMAChannel1 0 = DMAChannel0 12.5.5.2.16 USBDIntStatus

The USBDIntStatus register contains status bits that are related to conditions that can cause an interrupt to the CPU, if the corresponding interrupt enable bits are set in the USBDMask register. The field name extension Sticky implies that the status condition will remain registered until cleared by a CPU write of 1 to each bit of the field.

NOTE: There is no Ep0IrregPktSticky field because the default control EP will frequently receive packets that are not multiples of 32 bytes during normal operation.

TABLE 54 USBDIntStatus register format Write Field Name Bit(s) access Description CoreSuspendSticky 0 Clear Device core USB suspend flag. Sticky. 1 = USB suspend state. Set when device core udcvci_suspend signal transitions from 1 −> 0. 0 = default value. CoreUSBResetSticky 1 Clear Device core USB reset flag. Sticky. 1 = USB reset. Set when device core udcvci_reset signal transitions from 1 −> 0. 0 = default value. CoreUSBSOFSticky 2 Clear Device core USB Start Of Frame (SOF) flag. Sticky. 1 = USB SOF. Set when device core udcvci_sof signal transitions from 1 −> 0 0 = default value. CPUISITxBuffEmptySticky 3 Clear CPU ISI transmit buffer empty flag. Sticky. 1 = empty. 0 = default value. CPUEp0InBuffEmptySticky 4 Clear CPU EP0 IN buffer empty flag. Sticky. 1 = empty. 0 = default value. CPUEp5InBuffEmptySticky 5 Clear CPU EP5 IN buffer empty flag. Sticky. 1 = empty. 0 = default value. Ep0InNAKSticky 6 clear EP0-IN NAK flag. Sticky This flag is set if the USB device core issues a read request for EP0-IN and there is not a valid packet present in the EP0-IN buffer. The core will therefore send a NAK response to the IN token that was received from external USB host. This is an indicator of any back- pressure on the USB caused by EP0-IN. 1 = NAK sent. 0 = default value Ep5InNAKSticky 7 Clear Us per Ep0InNAK. Ep0OutNAKSticky 8 Clear EP0-OUT NAK flag. Sticky This flag is set if the USB device core issues a write request for EP0-OUT and there is no space in the OUT EP buffer for a the packet. The core will therefore send a NAK response to the OUT token that was received from external USB host. This is an indicator of any back-pressure on the USB caused by EP0- OUT. 1 = NAK sent. 0 = default value Ep1OutNAKSticky 9 Clear As per EP0OutNAK. Ep2OutNAKSticky 10 Clear As per EP0OutNAK. Ep3OutNAKSticky 11 Clear As per EP0OutNAK. Ep4OutNAKSticky 12 Clear As per EP0OutNAK. Ep1IrregPktSticky 13 Clear EP1-OUT irregular sized packet flag. Sticky. Indicates a packet that is not a multiple of 32 bytes in size was received by EP1-OUT. 1 = irregular sized packet received. 0 = default value. Ep2IrregPktSticky 14 Clear As per Ep1IrregPktSticky. Ep3IrregPktSticky 15 Clear As per Ep1IrregPktSticky. Ep4IrregPktSticky 16 Clear As per Ep1IrregPktSticky. OutBuffOverFlowSticky 17 Clear OUT EP buffer overflow flag. Sticky. This flag is set if the USB device core attempted to write a packet of more than 64 bytes to the OUT EP buffer. This is a fatal error, suggesting a problem in the USB device IP core. The SCB will take no further action. 1 = overflow condition detected. 0 = default value. InBuffUnderRunSticky 18 clear IN EP buffer underrun flag. Sticky. This flag is set if the USB device core attempted to read more data than was present from the IN EP buffer. This is a fatal error, suggesting a problem in the USB device IP core. The SCB will take no further action. 1 = underrun condition detected. 0 = default value. 12.5.5.2.17 USBDISIFIFOStatus

This register contains the status of the ISI mapped OUT EP packet FIFO. This is a secondary status register and will not cause any interrupts to the CPU.

TABLE 55 USBDISIFIFOStatus register format Write Field Name Bit(s) access Description Entry0Valid 0 none FIFO entry #0 valid field. This flag will be set by the USBD when the USB device core indicates the validity of packet entry #0 in the FIFO. 1 = valid USB packet in ISI OUT EP buffer 0. 0 = default value. Entry0Source 3:1 none FIFO entry #0 source field. Contains the EP associated with packet entry #0 in the FIFO. Binary Coded Decimal. Only valid when ISIBuff0PktValid = 1. Entry1Valid 4 none As per Entry0Valid. Entry1Source 7:5 none As per Entry0Source. Entry2Valid 8 none As per Entry0Valid. Entry2Source 11:9  none As per Entry0Source. Entry3Valid 12 none As per Entry0Valid. Entry3Source 15:13 none As per Entry0Source. 12.5.5.2.18 USBDDMAOFIFOStatus

This register description applies to USBDDMAOFIFOStatus and USBDDMA1FIFOStatus. This register contains the status of the DMAChannelN mapped OUT EP packet FIFO. This is a secondary status register and will not cause any interrupts to the CPU.

TABLE 56 USBDDMANFIFOStatus register format Write Field Name Bit(s) access Description Entry0Valid 0 none FIFO entry #0 valid field. This flag will be set by the USBD when the USB device core indicates the validity of packet entry #0 in the FIFO. 1 = valid USB packet in ISI OUT EP buffer 0. 0 = default value. Entry0Source 3:1 none FIFO entry #0 source field. Contains the EP associated with packet entry #0 in the FIFO. Binary Coded Decimal. Only valid when Entry0Valid = 1. Entry1Valid 4 none As per Entry0Valid. Entry1Source 7:5 none As per Entry0Source. 12.5.5.2.19 USBDResume

This register causes the USB device core to initiate resume signalling to the external USB host. Only applicable when the device core is in the suspend state.

TABLE 57 USBDResume register format Write Field Name Bit(s) access Description USBDResume 0 full USBD core resume register. The USBD will clear this register upon resume notification from the device core. 1 = generate resume signalling. 0 = default value. 12.5.5.2.20 USBDSetup

This register controls the general setup/configuration of the USBD.

TABLE 58 USBDSetup register format write Field Name Bit(s) access Description Ep1IrregPktCntrl 0 full EP 1 OUT irregular sized packet control. An Irregular sized packet is defined as a packet that is not a multiple of 32 bytes. 1 = discard irregular sized packets. 0 = read 32 bytes from buffer, regardless of packets size. Ep2IrregPktCntrl 1 full As per Ep1IrregPktDiscard Ep3IrregPktCntrl 2 full As per Ep1IrregPktDiscard Ep4IrregPktCntrl 3 full As per Ep1IrregPktDiscard 12.5.5.2.21 USBDEpNInBuffCtrl Register

This register description applies to USBDEp0InBuffCtrl and USBDEp5InBuffCtrl.

TABLE 59 USBDEpNInBuffCtrl register format Write Field Name Bit(s) access Description PktValid 0 full Setting this register validates the contents of USBDEpNInBuff. This field will be cleared by the SCB once the packet has been successfully transmitted to the external USB host. NOTE: The CPU should not clear this field under normal operation. If the CPU clears this field during a packet transfer to the USB, the transfer will be aborted - this is not recommended. 1 = valid packet. 0 = default state. 12.5.5.2.22 USBDMask

This register serves as an interrupt mask for all USBD status conditions that can cause a CPU interrupt. Setting a field enables interrupt generation for the associated status event. Clearing a field disables interrupt generation for the associated status event. All interrupts will be generated in an edge sensitive manner, i.e. when the associated status register transitions from 0->1.

TABLE 60 USBDMask register format Write Field Name Bit(s) access Description CoreSuspendStickyEn 0 full CoreSuspendSticky status interrupt enable. CoreUSBResetStickyEn 1 full CoreUSBResetSticky status interrupt enable. CoreUSBSOFStickyEn 2 full CoreUSBSOFSticky status interrupt enable. CPUISITxBuffEmptyStickyEn 3 full CPUISITxBuffEmptySticky status interrupt enable. CPUEp0InBuffEmptyStickyEn 4 full CPUEp0InBuffEmptySticky status interrupt enable. CPUEp5InBuffEmptyStickyEn 5 full CPUEp5InBuffEmptySticky status interrupt enable. Ep0InNAKStickyEn 6 full Ep0InNAKSticky status interrupt enable. Ep5InNAKStickyEn 7 full Ep5InNAKSticky status interrupt enable. Ep0OutNAKStickyEn 8 full Ep0OutNAKSticky status interrupt enable. Ep1OutNAKStickyEn 9 full Ep1OutNAKSticky status interrupt enable. Ep2OutNAKStickyEn 10 full Ep2OutNAKSticky status interrupt enable. Ep3OutNAKStickyEn 11 full Ep3OutNAKSticky status interrupt enable. Ep4OutNAKStickyEn 12 full Ep4OutNAKSticky status interrupt enable. Ep1IrregPktStickyEn 13 full Ep1IrregPktSticky status interrupt enable. Ep2IrregPktStickyEn 14 full Ep2IrregPktSticky status interrupt enable. Ep3IrregPktStickyEn 15 full Ep3IrregPktSticky status interrupt enable. Ep4IrregPktStickyEn 16 full Ep4IrregPktSticky status interrupt enable. OutBuffOverFlowStickyEn 17 full OutBuffOverFlowSticky status interrupt enable. InBuffUnderRunStickyEn 18 full InBuffUnderRunSticky status interrupt enable. 12.5.5.2.23 USBDDebug

This register is intended for debug purposes only. Contains non-sticky versions of all interrupt capable status bits, which are referred to as dynamic in the table.

TABLE 61 USBDDebug register format write Field Name Bit(s) access Description CoreTimeStamp 10:0 none USB device core frame number. CoreSuspend 11 none Dynamic version of CoreSuspendSticky. CoreUSBReset 12 none Dynamic version of CoreUSBResetSticky. CoreUSBSOF 13 none Dynamic version of CoreUSBSOFSticky. CPUISITxBuffEmpty 14 none Dynamic version of CPUISITxBuffEmptySticky. CPUEp0InBuffEmpty 15 none Dynamic version of CPUEp0InBuffEmptySticky. CPUEp5InBuffEmpty 16 none Dynamic version of CPUEp5InBuffEmptySticky. Ep0InNAK 17 none Dynamic version of Ep0InNAKSticky. Ep5InNAK 18 none Dynamic version of Ep5InNAKSticky. Ep0OutNAK 19 none Dynamic version of Ep0OutNAKSticky. Ep1OutNAK 20 none Dynamic version of Ep1OutNAKSticky. Ep2OutNAK 21 none Dynamic version of Ep2OutNAKSticky. Ep3OutNAK 22 none Dynamic version of Ep3OutNAKSticky. Ep4OutNAK 23 none Dynamic version of Ep4OutNAKSticky. Ep1IrregPkt 24 none Dynamic version of Ep1IrregPktSticky. Ep2IrregPkt 25 none Dynamic version of Ep2IrregPktSticky. Ep3IrregPkt 26 none Dynamic version of Ep3IrregPktSticky. Ep4IrregPkt 27 none Dynamic version of Ep4IrregPktSticky. OutBuffOverFlow 28 none Dynamic version of OutBuffOverFlowSticky. InBuffUnderRun 29 none Dynamic version of InBuffUnderRunSticky. 12.5.5.2.24 USBHStatus

This register contains all status bits associated with the USBH. The field name extension Sticky implies that the status condition will remain registered until cleared by a CPU write.

TABLE 62 USBHStatus register format Write Field Name Bit(s) access Description CoreIRQSticky 0 clear HC core IRQ interrupt flag. Sticky Set when HC core UHOSTC_IrqN output signal transitions from 0 −> 1. Refer to OHCI spec for details on HC interrupt processing. 1 = IRQ interrupt from core. 0 = default value. CoreSMISticky 1 clear HC core SMI interrupt flag. Sticky Set when HC core UHOSTC_SmiN output signal transi- tions from 0 −> 1. Refer to OHCI spec for details on HC interrupt processing. 1 = SMI interrupt from HC. 0 = default value. CoreBuffAcc 2 none HC core buffer access flag. HC core UHOSTC_BufAcc output signal. Indicates whether the HC is accessing a descriptor or a buffer in shared system memory. 1 = buffer access 0 = descriptor access. 12.5.5.2.25 USBHMask

This register serves as an interrupt mask for all USBH status conditions that can cause a CPU interrupt. All interrupts will be generated in an edge sensitive manner, i.e. when the associated status register transitions from 0->1.

TABLE 63 USBHMask register format Write Field Name Bit(s) access Description CoreIRQIntEn 0 full CoreIRQSticky status interrupt enable. 1 = enable. 0 = disable. CoreSMIIntEn 1 full CoreSMISticky status interrupt enable. 1 = enable. 0 = disable. 12.5.5.2.26 USBHDebug

This register is intended for debug purposes only. Contains non-sticky versions of all interrupt capable status bits, which are referred to as dynamic in the table.

TABLE 64 USBHDebug register format write Field Name Bit(s) access Description CoreIRQ 0 none Dynamic version of CoreIRQSticky. CoreSMI 1 None Dynamic version of CoreSMISticky. 12.5.5.2.27 ISICntrl

This register controls the general setup/configuration of the ISI.

Note that the reset value of this register allows the SoPEC to automatically become an ISIMaster (AutoMasterEnable=1) if any USB packets are received on endpoints 24. On becoming an ISIMaster the ISIMasterSel bit is set and any USB or CPU packets destined for other ISI devices are transmitted. The CPU can override this capability at any time by clearing the AutoMasterEnable bit.

TABLE 65 ISICntrl register format Write Field Name Bit(s) access Description TxEnable 0 Full ISI transmit enable. Enables ISI transmission of long or ping packets. ACKs may still be transmitted when this bit is 0. This is cleared by transmit errors and needs to be restarted by the CPU. 1 = Transmission enabled 0 = Transmission disabled RxEnable 1 Full ISI receive enable. Enables ISI reception. This is can only be cleared by the CPU and it is only anticipated that reception will be disabled when the ISI in not in use and the ISI pins are being used by the GPIO for another purpose. 1 = Reception enabled 0 = Reception disabled ISIMasterSel 2 Full ISI master select. Determines whether the SoPEC is an ISIMaster or not 1 = ISIMaster 0 = ISISlave AutoMasterEnable 3 Full ISI auto master enable. Enables the device to automatically become the ISIMaster if activity is detected on USB endpoints 2-4. 1 = auto-master operation enabled 0 = auto-master operation disabled 12.5.5.2.28 ISIId

TABLE 66 ISIId register format Field Write Name Bit(s) access Description ISIId 3:0 Full ISIId for this SoPEC. SoPEC resets to being an ISISlave with ISIId0. 0xF (the broadcast ISIId) is an illegal value and should not be written to this register. 12.5.5.2.29 ISINumRetries

TABLE 67 ISINumRetries register format Write Field Name Bit(s) access Description ISINumRetries 3:0 Full Number of ISI retransmissions to attempt in response to an inferred NAK before aborting a long packet transmission 12.5.5.2.30 ISIPingScheduleN

This register description applies to ISIPingSchedule0, ISIPingSchedule1 and ISIPingSchedule2.

TABLE 68 ISIPingScheduleN register format Write Field Name Bit(s) access Description ISIPingSchedule 14:0 Full Denotes which ISIIds will be receive ping packets. Note that bit0 refers to ISIId0, bit1 to ISIId1 . . . bit14 to ISIId14. 12.5.5.2.31 ISITotalPeriod

TABLE 69 ISITotalPeriod register format Write Field Name Bit(s) access Description ISITotalPeriod 3:0 Full Reload value of the ISITotalPeriod counter 12.5.5.2.32 ISILocalPeriod

TABLE 70 ISILocalPeriod register format Write Field Name Bit(s) access Description ISILocalPeriod 3:0 Full Reload value of the ISILocalPeriod counter 12.5.5.2.33 ISIIntStatus

The ISIIntStatus register contains status bits that are related to conditions that can cause an interrupt to the CPU, if the corresponding interrupt enable bits are set in the ISIMask register.

TABLE 71 ISIIntStatus register Write Field Name Bit(s) access Description TxErrorSticky 0 None ISI transmit error flag. Sticky. Receiving ISI device would not accept the transmitted packet. Only set after NumRetries unsuccessful retransmissions. (excluding ping packets). This bit is cleared by the ISI after transmission has been re-enabled by the CPU setting the TxEnable bit of the ISICntrl register. 1 = transmit error. 0 = default state. RxFrameErrorSticky 1 Clear ISI receive framing error flag. Sticky. This bit is set by the ISI when a framing error detect- ed in the received packet, which can be caused by an incorrect Start or Stop field or by bit stuffing errors. 1 = framing error detected. 0 = default state. RxCRCErrorSticky 2 Clear ISI receive CRC error flag. This bit is set by the ISI when a CRC error is detected in an incoming packet. Other than dropping the errored packet ISI reception is un- affected by a CRC Error. 1 = CRC error 0 = default state. RxBuffOverFlowSticky 3 Clear ISI receive buffer over flow flag. Sticky. An overflow has occurred in the ISI receive buffer and a packet had to be dropped. 1 = over flow condition detected. 0 = default state. 12.5.5.2.34 ISITxBuffStatus

The ISITxBuffStatus register contains status bits that are related to the ISI Tx buffer. This is a secondary status register and will not cause any interrupts to the CPU.

TABLE 72 ISITxBuffStatus register format Write Field Name Bit(s) access Description Entry0PktValid 0 None ISI Tx buffer entry #0 packet valid flag. This flag will be set by the ISI when a valid ISI packet is written to entry #0 in the ISITxBuff for transmission over the ISI bus. A Tx packet is considered valid when it is 32 bytes in size and the ISI has written the packet header information to Entry0PktDesc, Entry0DestISIId and Entry0DestISISubId. 1 = packet valid. 0 = default value. Entry0PktDesc 3:1 None ISI Tx buffer entry #0 packet descriptor. PktDesc field as per Table for the packet entry #0 in the ISITxBuff. Only valid when Entry0PktValid = 1. Entry0DestISIId 7:4 None ISI Tx buffer entry #0 destination ISI ID. Denotes the ISIId of the target SoPEC as per Table . Only valid when Entry0PktValid = 1. Entry0DestISISubId 8 None ISI Tx buffer entry #0 destination ISI sub ID. Indicates which DMAChannel on the target SoPEC that packet entry #0 in the ISITxBuff is destined for. Only valid when Entry0PktValid = 1. 1 = DMAChannel1 0 = DMAChannel0 Entry1PktValid 9 None As per Entry0PktValid. Entry1PktDesc 12:10 None As per Entry0PktDesc. Entry1DestISIId 16:13 None As per Entry0DestISIId. Entry1DestISISubId 17 None As per Entry0DestISISubId. 12.5.5.2.35 ISIRxBuffStatus

The ISIRxBuffStatus register contains status bits that are related to the ISI Rx buffer. This is a secondary status register and will not cause any interrupts to the CPU.

TABLE 73 ISIRxBuffStatus register format Write Field Name Bit(s) access Description Entry0PktValid 0 None ISI Rx buffer entry #0 packet valid flag. This flag will be set by the ISI when a valid ISI packet is received and written to entry #0 of the ISIRxBuff. A Rx packet is considered valid when it is 32 bytes in size and no framing or CRC errors were detected. 1 = valid packet 0 = default value Entry0PktDesc 3:1 None ISI Rx buffer entry #0 packet descriptor. PktDesc field as per Table for packet entry #0 of the ISIRxBuff. Only valid when Entry0PktValid = 1. Entry0DestISIId 7:4 None ISI Rx buffer 0 destination ISI ID. Denotes the ISIId of the target SoPEC as per Table . This should always corre- spond to the local SoPEC ISIId. Only valid when Entry0PktValid = 1. Entry0DestISISubId 8 None ISI Rx buffer 0 destination ISI sub ID. Indicates which DMAChannel on the target SoPEC that entry #0 of the ISIRxBuff is destined for. Only valid when Entry0PktValid = 1. 1 = DMAChannel1 0 = DMAChannel0 Entry1PktValid 9 None As per Entry0PktValid. Entry1PktDesc 12:10 None As per Entry0PktDesc. Entry1DestISIId 16:13 None As per Entry0DestISIId. Entry1DestISISubId 17 None As per Entry0DestISISubId. 12.5.5.2.36 ISIMask Register

An interrupt will be generated in an edge sensitive manner i.e. the ISI will generate an isi_icu_irq pulse each time a status bit goes high and the corresponding bit of the ISIMask register is enabled.

TABLE 74 ISIMask register Write Field Name Bit(s) access Description TxErrorIntEn 0 Full TxErrorSticky status interrupt enable. 1 = enable. 0 = disable. RxFrameErrorIntEn 1 Full RxFrameErrorSticky status interrupt enable. 1 = enable. 0 = disable. RxCRCErrorIntEn 2 Full RxCRCErrorSticky status interrupt enable. 1 = enable. 0 = disable. RxBuffOverFlowIntEn 3 Full RxBuffOverFlowSticky status interrupt enable. 1 = enable. 0 = disable. 12.5.5.2.37 ISISubIdNSeq

This register description applies to ISISubId0Seq and ISISubId0Seq.

TABLE 75 ISISubIdNSeq register format Write Field Name Bit(s) access Description ISISubIdNSeq 0 Full ISI sub ID channel N sequence bit. This bit may be initialised by the CPU but is updated by the ISI each time an error-free long packet is recieved. 12.5.5.2.38 ISISubIdSeqMask

TABLE 76 ISISubIdSeqMask register format Write Field Name Bit(s) access Description ISISubIdSeq0Mask 0 Full ISI sub ID channel 0 sequence bit mask. Setting this bit ensures that the sequence bit will be ignored for incoming packets for the ISISubId. 1 = ignore sequence bit. 0 = default state. ISISubIdSeq1Mask 1 Full As per ISISubIdSeq0Mask. 12.5.5.2.39 ISINumPins

TABLE 77 ISINumPins register format Write Field Name Bit(s) access Description ISINumPins 0 Full Select number of active ISI pins. 1 = 4 pins 0 = 2 pins 12.5.5.2.40 ISITurnAround

The ISI bus turnaround time will reset to its maximum value of 0xF to provide a safer starting mode for the ISI bus. This value should be set to a value that is suitable for the physical implementation of the ISI bus, i.e. the lowest turn around time that the physical implementation will allow without significant degradation of signal integrity.

TABLE 78 ISITurnAround register format Write Field Name Bit(s) access Description ISITurnAround 3:0 Full ISI bus turn around time in ISI clock cycles (32 MHz). 12.5.5.2.41 ISIShortReplyWin

The ISI short packet reply window time will reset to its maximum value of 0x1F to provide a safer starting mode for the ISI bus. This value should be set to a value that will allow for expected frequency of bit stuffing and receiver response timing.

TABLE 79 ISIShortReplyWin register format Write Field Name Bit(s) access Description ISIShortReplyWin 4:0 Full ISI long packet reply window in ISI clock cycles (32 MHz). 12.5.5.2.42 ISILongReplyWin

The ISI long packet reply window time will reset to its maximum value of 0x1FF to provide a safer starting mode for the ISI bus. This value should be set to a value that will allow for expected frequency of bit stuffing and receiver response timing.

TABLE 80 ISILongReplyWin register format Write Field Name Bit(s) access Description ISILongReplyWin 8:0 Full ISI long packet reply window in ISI clock cycles (32 MHz). 12.5.5.2.43 ISIDebug

This register is intended for debug purposes only. Contains non-sticky versions of all interrupt capable status bits, which are referred to as dynamic in the table.

TABLE 81 ISIDebug register format Write Field Name Bit(s) access Description TxError 0 None Dynamic version of TxErrorSticky. RxFrameError 1 None Dynamic version of RxFrameErrorSticky. RxCRCError 2 None Dynamic version of RxCRCErrorSticky. RxBuffOverFlow 3 None Dynamic version of RxBuffOverFlowSticky. 12.5.5.3 CPU Bus Interface 12.5.5.4 Control Core Logic 12.5.5.5 DIU Bus Interface 12.6 DMA Regs

All of the circular buffer registers are 256-bit word aligned as required by the DIU. The DMAnBottomAdr and DMAnTopAdr registers are inclusive i.e. the addresses contained in those registers form part of the circular buffer. The DMAnCurrWPtr always points to the next location the DMA manager will write to so interrupts are generated whenever the DMA manager reaches the address in either the DMAnIntAdr or DMAnMaxAdr registers rather than when it actually writes to these locations. It therefore can not write to the location in the DMAnMaxAdr register.

SCB Map Regs

The SCB map is configured by mapping a USB endpoint on to a data sink. This is performed on a endpoint basis i.e. each endpoint has a configuration register to allow its data sink be selected. Mapping an endpoint on to a data sink does not initiate any data flow—each endpoint/data sink needs to be enabled by writing to the appropriate configuration registers for the USBD, ISI and DMA manager.

13. General Purpose IO (GPIO)

13.1 Overview

The General Purpose IO block (GPIO) is responsible for control and interfacing of GPIO pins to the rest of the SoPEC system. It provides easily programmable control logic to simplify control of GPIO functions. In all there are 32 GPIO pins of which any pin can assume any output or input function. Possible output functions are

-   -   4 Stepper Motor control Outputs     -   12 Brushless DC Motor Control Output (total of 2 different         controllers each with 6 outputs)     -   4 General purpose high drive pulsed outputs capable of driving         LEDs.     -   4 Open drain IOs used for LSS interfaces     -   4 Normal drive low impedance IOs used for the ISI interface in         Multi-SoPEC mode

Each of the pins can be configured in either input or output mode, each pin is independently controlled. A programmable de-glitching circuit exists for a fixed number of input pins. Each input is a schmidt trigger to increase noise immunity should the input be used without the de-glitch circuit. The mapping of the above functions and their alternate use in a slave SoPEC to GPIO pins is shown in Table 82 below.

TABLE 82 GPIO pin type GPIO pin(s) Pin IO Type Default Function gpio[3:0] Normal drive, low Pins 1 and 0 in ISI Mode, impedance IO (35 pins 2 and 3 in input mode Ohm), Integrated pull-up resistor gpio[7:4] High drive, normal Input Mode impedance IO (65 Ohm), intended for LED drivers gpio[31:8] Normal drive, Input Mode normal impedance IO (65 Ohm), no pull-up 13.2 Stepper Motor Control

The motor control pins can be directly controlled by the CPU or the motor control logic can be used to generate the phase pulses for the stepper motors. The controller consists of two central counters from which the control pins are derived. The central counters have several registers (see Table) used to configure the cycle period, the phase, the duty cycle, and counter granularity.

There are two motor master counters (0 and 1) with identical features. The period of the master counters are defined by the MotorMasterClkPeriod[1:0] and MotorMasterClkSrc registers i.e. both master counters are derived from the same MotorMasterClkSrc. The MotorMasterClkSrc defines the timing pulses used by the master counters to determine the timing period. The MotorMasterClkSrc can select clock sources of 1 μs, 100 μs, 10 ms and pclk timing pulses.

The MotorMasterClkPeriod[1:0] registers are set to the number of timing pulses required before the timing period re-starts. Each master counter is set to the relevant MotorMasterClkPeriod value and counts down a unit each time a timing pulse is received.

The master counters reset to MotorMasterClkPeriod value and count down. Once the value hits zero a new value is reloaded from the MotorMasterClkPeriod[1:0] registers. This ensures that no master clock glitch is generated when changing the clock period.

Each of the IO pins for the motor controller are derived from the master counters. Each pin has independent configuration registers. The MotorMasterClkSelect[3:0] registers define which of the two master counters to use as the source for each motor control pin. The master counter value is compared with the configured MotorCtrlLow and MotorCtrlHigh registers (bit fields of the MotorCtrlConfig register). If the count is equal to MotorCtrlHigh value the motor control is set to 1, if the count is equal to MotorCtrlLow value the motor control pin is set to 0.

This allows the phase and duty cycle of the motor control pins to be varied at pclk granularity. The motor control generators keep a working copy of the MotorCtrlLow, MotorCtrlHigh values and update the configured value to the working copy when it is safe to do so. This allows the phase or duty cycle of a motor control pin to be safely adjusted by the CPU without causing a glitch on the output pin.

Note that when reprogramming the MotorCtrlLow, MotorCtrlHigh registers to reorder the sequence of the transition points (e.g changing from low point less than high point to low point greater than high point and vice versa) care must still taken to avoid introducing glitching on the output pin.

13.3 LED Control

LED lifetime and brightness can be improved and power consumption reduced by driving the LEDs with a pulsed rather than a DC signal. The source clock for each of the LED pins is a 7.8 kHz (128 μs period) clock generated from the 1 μs clock pulse from the Timers block. The LEDDutySelect registers are used to create a signal with the desired waveform. Unpulsed operation of the LED pins can be achieved by using CPU IO direct control, or setting LEDDutySelect to 0. By default the LED pins are controlled by the LED control logic.

13.4 LSS Interface via GPIO

In some SoPEC system configurations one or more of the LSS interfaces may not be used. Unused LSS interface pins can be reused as general IO pins by configuring the IOModeSelect registers. When a mode select register for a particular GPIO pin is set to 23,22,21,20 the GPIO pin is connected to LSS control IOs 3 to 0 respectively.

13.5 ISI Interface via GPIO

In Multi-SoPEC mode the SCB block (in particular the ISI sub-block) requires direct access to and from the GPIO pins. Control of the ISI interface pins is determined by the IOModeSelect registers. When a mode select register for a particular GPIO pin is set to 27,26,25,24 the GPIO pin connected to the ISI control bits 3 to 0 respectively. By default the GPIO pins 1 to 0 are directly controlled by the ISI block.

In single SoPEC systems the pins can be re-used by the GPIO.

13.6 CPU GPIO Control

The CPU can assume direct control of any (or all) of the IO pins individually. On a per pin basis the CPU can turn on direct access to the pin by configuring the IOModeSelect register to CPU direct mode. Once set the IO pin assumes the direction specified by the CpuIODirection register. When in output mode the value in register CpuIOOut will be directly reflected to the output driver. When in input mode the status of the input pin can be read by reading CpuIOIn register. When writing to the CpuIOOut register the value being written is XORed with the current value in CpuIOOut. The CPU can also read the status of the 10 selected de-glitched inputs by reading the CpuIOInDeGlitch register.

13.7 Programmable De-Glitching Logic

Each IO pin can be filtered through a de-glitching logic circuit, the pin that the de-glitching logic is connected to is configured by the InputPinSelect registers. There are IO de-glitching circuits, so a maximum of IO input pin can be de-glitched at anytime.

The de-glitch circuit can be configured to sample the IO pin for a predetermined time before concluding that a pin is in a particular state. The exact sampling length is configurable, but each de-glitch circuit must use one of two possible configured values (selected by DeGlitchSelect). The sampling length is the same for both high and low states. The DeGlitchCount is programmed to the number of system time units that a state must be valid for before the state is passed on. The time units are selected by DeGlitchClkSel and can be one of 1 μs, 100 μs, 10 ms and pclk pulses.

For example if DeGlitchCount is set to 10 and DeGlitchClkSel set to 3, then the selected input pin must consistently retain its value for 10 system clock cycles (pclk) before the input state will be propagated from CpuIOIn to CpuIOInDeglitch.

13.8 Interrupt Generation

Any of the selected input pins (selected by InputPinSelect) can generate an interrupt from the raw or deglitched version of the input pin. There are IO possible interrupt sources from the GPIO to the interrupt controller, one interrupt per input pin. The InterruptSrcSelect register determines whether the raw input or the deglitched version is used as the interrupt source.

The interrupt type, masking and priority can be programmed in the interrupt controller.

13.9 Frequency Analyser

The frequency analyser measures the duration between successive positive edges on a selected input pin (selected by InputPinSelect) and reports the last period measured (FreqAnaLastPeriod) and a running average period (FreqAnaAverage).

The running average is updated each time a new positive edge is detected and is calculated by FreqAnaAverage=(FreqAnaAverage/8)*7+FreqAnaLastPeriod/8.

The analyser can be used with any selected input pin (or its deglitched form), but only one input at a time can be selected. The input is selected by the FreqAnaPinSelect (range of 0 to 9) and its deglitched form can be selected by FreqAnaPinFormSelect.

13.10 Brushless DC (BLDC) Motor Controllers

The GPIO contains 2 brushless DC (BLDC) motor controllers. Each controller consists of 3 hall inputs, a direction input, and six possible outputs. The outputs are derived from the input state and a pulse width modulated (PWM) input from the Stepper Motor controller, and is given by the truth table in Table 83.

TABLE 83 Truth Table for BLDC Motor Controllers direction hc hb ha q6 q5 q4 q3 q2 q1 0 0 0 1 0 0 0 1 PWM 0 0 0 1 1 PWM 0 0 1 0 0 0 0 1 0 PWM 0 0 0 0 1 0 1 1 0 0 0 PWM 0 0 1 0 1 0 0 0 1 PWM 0 0 0 0 1 0 1 0 1 0 0 PWM 0 0 0 0 0 0 0 0 0 0 0 0 1 1 1 0 0 0 0 0 0 1 0 0 1 0 0 PWM 0 0 1 1 0 1 1 PWM 0 0 0 0 1 1 0 1 0 PWM 0 0 1 0 0 1 1 1 0 0 0 0 1 PWM 0 1 1 0 0 0 1 0 0 PWM 0 1 1 0 1 0 1 PWM 0 0 0 1 0 0 0 0 0 0 0 0 0 1 1 1 1 0 0 0 0 0 0

All inputs to a BLDC controller must be de-glitched. Each controller has its inputs hardwired to de-glitch circuits. Controller 1 hall inputs are de-glitched by circuits 2 to 0, and its direction input is de-glitched by circuit 3. Controller 2 inputs are de-glitched by circuits 6 to 4 for hall inputs and 7 for direction input.

Each controller also requires a PWM input. The stepper motor controller outputs are reused, output 0 is connected to BLDC controller 1, and output 1 to BLDC controller 2.

The controllers have two modes of operation, internal and external direction control (configured by BLDCMode). If a controller is in external direction mode the direction input is taken from a de-glitched circuit, if it is in internal direction mode the direction input is configured by the BLDCDirection register.

The BLDC controller outputs are connected to the GPIO output pins by configuring the IOModeSelect register for each pin. e.g Setting the mode register to 8 will connect q1 Controller 1 to drive the pin.

13.11 Implementation

13.11.1 Definitions of I/O

TABLE 84 I/O definition Port name Pins I/O Description Clocks and Resets Pclk 1 In System Clock prst_n 1 In System reset, synchronous active low tim_pulse[2:0] 3 In Timers block generated timing pulses. 0 - 1 μs pulse 1 - 100 μs pulse 2 - 10 ms pulse CPU Interface cpu_adr[8:2] 8 In CPU address bus. Only 7 bits are required to decode the address space for this block cpu_dataout[31:0] 32 In Shared write data bus from the CPU gpio_cpu_data[31:0] 32 Out Read data bus to the CPU cpu_rwn 1 In Common read/not-write signal from the CPU cpu_gpio_sel 1 In Block select from the CPU. When cpu_gpio_sel is high both cpu_adr and cpu_dataout are valid gpio_cpu_rdy 1 Out Ready signal to the CPU. When gpio_cpu_rdy is high it indicates the last cycle of the access. For a write cycle this means cpu_dataout has been registered by the GPIO block and for a read cycle this means the data on gpio_cpu_data is valid. gpio_cpu_berr 1 Out Bus error signal to the CPU indicating an invalid access. gpio_cpu_debug_valid 1 Out Debug Data valid on gpio_cpu_data bus. Active high cpu_acode[1:0] 2 In CPU Access Code signals. These decode as follows: 00 - User program access 01 - User data access 10 - Supervisor program access 11 - Supervisor data access IO Pins gpio_o[31:0] 32 Out General purpose IO output to IO driver gpio_i[31:0] 32 In General purpose IO input from IO receiver gpio_e[31:0] 32 Out General purpose IO output control. Active high driving GPIO to LSS lss_gpio_dout[1:0] 2 In LSS bus data output Bit 0 - LSS bus 0 Bit 1 - LSS bus 1 gpio_lss_din[1:0] 2 Out LSS bus data input Bit 0 - LSS bus 0 Bit 1 - LSS bus 1 lss_gpio_e[1:0] 2 In LSS bus data output enable, active high Bit 0 - LSS bus 0 Bit 1 - LSS bus 1 lss_gpio_clk[1:0] 2 In LSS bus clock output Bit 0 - LSS bus 0 Bit 1 - LSS bus 1 GPIO to ISI gpio_isi_din[1:0] 2 Out Input data from IO receivers to ISI. isi_gpio_dout[1:0] 2 In Data output from ISI to IO drivers isi_gpio_e[1:0] 2 In GPIO ISI pins output enable (active high) from ISI interface usbh_gpio_power_en 1 In Port Power enable from the USB host core, active high gpio_usbh_over_current 1 Out Over current detect to the USB host core, active high Miscellaneous gpio_icu_irq[9:0] 10 Out GPIO pin interrupts gpio_cpr_wakeup 1 Out SoPEC wakeup to the CPR block active high. Debug debug_data_out[31:0] 32 In Output debug data to be muxed on to the GPIO pins debug_cntrl[31:0] 32 In Control signal for each GPIO bound debug data line indicating whether or not the debug data should be selected by the pin mux 13.11.2 Configuration Registers

The configuration registers in the GPIO are programmed via the CPU interface. Refer to section 11.4.3 on page 69 for a description of the protocol and timing diagrams for reading and writing registers in the GPIO. Note that since addresses in SoPEC are byte aligned and the CPU only supports 32-bit register reads and writes, the lower 2 bits of the CPU address bus are not required to decode the address space for the GPIO. When reading a register that is less than 32 bits wide zeros should be returned on the upper unused bit(s) of gpio_cpu_data. Table 85 lists the configuration registers in the GPIO block

TABLE 85 GPIO Register Definition Address GPIO_base+ Register #bits Reset Description 0x000-0x07C IOModeSelect[31:0] 32 × 5 See Specifies the mode of operation for each Table for GPIO pin. One 5 bit bus per pin. default values Possible assignment values and correspond controller outputs are as follows Value - Controlled by 3 to 0 - Output, LED controller 4 to 1 7 to 4 - Output Stepper Motor control 4-1 13 to 8 - Output BLDC 1 Motor control 6-1 19 to 14 - Output BLDC 2 Motor control 6-1 23 to 20 - LSS control 4-1 27 to 24 - ISI control 4-1 28 - CPU Direct Control 29 - USB power enable output 30 - Input Mode 0x080-0xA4 InputPinSelect[9:0] 10 × 5 0x00 Specifies which pins should be selected as inputs. Used to select the pin source to the DeGlitch Circuits. CPU IO Control 0x0B0 CpuIOUserModeMask 32 0x0000_0000 User Mode Access Mask to CPU GPIO control register. When 1 user access is enabled. One bit per gpio pin. Enables access to CpuIODirection, CpuIOOut and CpuIOIn in user mode. 0x0B4 CpuIOSuperModeMask 32 0xFFFF_FFFF Supervisor Mode Access Mask to CPU GPIO control register. When 1 supervisor access is enabled. One bit per gpio pin. Enables access to CpuIODirection, CpuIOOut and CpuIOIn in supervisor mode. 0x0B8 CpuIODirection 32 0x0000_0000 Indicates the direction of each IO pin, when controlled by the CPU 0 - Indicates Input Mode 1 - Indicates Output Mode 0x0BC CpuIOOut 32 0x0000_0000 Value used to drive output pin in CPU direct mode. bits31:0 - Value to drive on output GPIO pins When written to the register assumes the new value XORed with the current value. 0x0C0 CpuIOIn 32 External Value received on each input pin regardless pin value of mode. Read Only register. 0x0C4 CpuDeGlitchUserModeMask 10 0x000 User Mode Access Mask to CpuIOInDeglitch control register. When 1 user access is enabled, otherwise bit reads as zero. 0x0C8 CpuIOInDeglitch 10 0x000 Deglitched version of selected input pins. The input pins are selected by the InputPinSelect register. Note that after reset this register will reflect the external pin values 256 pclk cycles after they have stabilized. Read Only register. Deglitch control 0x0D0-0x0D4 DeGlitchCount[1:0] 2 × 8 0xFF Deglitch circuit sample count in DeGlitchClkSrc selected units. 0x0D8-0x0DC DeGlitchClkSrc[1:0] 2 × 2 0x3 Specifies the unit use of the GPIO deglitch circuits: 0 - 1 μs pulse 1 - 100 μs pulse 2 - 10 ms pulse 3 - pclk 0x0E0 DeGlitchSelect 10 0x000 Specifies which deglitch count (DeGlitchCount) and unit select (DeGlitchClkSrc) should be used with each de-glitch circuit 0 - Specifies DeGlitchCount[0] and DeGlitchClkSrc[0] 1 - Specifies DeGlitchCount[1] and DeGlitchClkSrc[1] Motor Control 0x0E4 MotorCtrlUserModeEnable 1 0x0 User Mode Access enable to Motor control configuration registers. When 1 user access is enabled. Enables user access to MotorMasterClkPeriod, MotorMasterClkSrc, MotorDutySelect, MotorPhaseSelect, MotorMasterClockEnable, MotorMasterClkSelect, BLDCMode and BLDCDirection registers 0x0E8-0x0EC MotorMasterClkPeriod[1:0] 2 × 16 0x0000 Specifies the motor controller master clock periods in MotorMasterClkSrc selected units 0x0F0 MotorMasterClkSrc 2 0x0 Specifies the unit use by the motor controller master clock generator: 0 - 1 μs pulse 1 - 100 μs pulse 2 - 10 ms pulse 3 - pclk 0x0F4-0x100 MotorCtrlConfig[3:0] 4 × 32 0x0000_0000 Specifies the transition points in the clock period for each motor control pin. One register per pin bits 15:0 - MotorCtrlLow, high to low transition point bits 31:16 - MotorCtrlHigh, low to high transition point 0x104 MotorMasterClkSelect 4 0x0 Specifies which motor master clock should be used as a pin generator source 0 - Clock derived from MotorMasterClockPeriod[0] 1 -Clock derived from MotorMasterClockPeriod[1] 0x108 MotorMasterClockEnable 2 0x0 Enable the motor master clock counter. When 1 count is enabled Bit 0 - Enable motor master clock 0 Bit 1 - Enable motor master clock 1 BLDC Motor Controllers 0x10C BLDCMode 2 0x0 Specifies the Mode of operation of the BLDC Controller. One bit per Controller. 0 - External direction control 1 - Internal direction control 0x110 BLDCDirection 2 0x0 Specifies the direction input of the BLDC controller. Only used when BLDC controller is an internal direction control mode. One bit per controller. LED control 0x114 LEDCtrlUserModeEnable 4 0x0 User Mode Access enable to LED control configuration registers. When 1 user access is enabled. One bit per LEDDutySelect select register. 0x118-0x124 LEDDutySelect3:0] 4 × 3 0x0 Specifies the duty cycle for each LED control output. See FIG. 54 for encoding details. The LEDDutySelect[3:0] registers determine the duty cycle of the LED controller outputs Frequency Analyser 0x130 FreqAnaUserModeEnable 1 0x0 User Mode Access enable to Frequency analyser configuration registers. When 1 user access is enabled. Controls access to FreqAnaPinFormSelect, FreqAnaLastPeriod, FreqAnaAverage and FreqAnaCountInc. 0x134 FreqAnaPinSelect 4 0x00 Selects which selected input should be used for the frequency analyses. 0x138 FreqAnaPinFormSelect 1 0x0 Selects if the frequency analyser should use the raw input or the deglitched form. 0 - Deglitched form of input pin 1 - Raw form of input pin 0x13C FreqAnaLastPeriod 16 0x0000 Frequency Analyser last period of selected input pin. 0x140 FreqAnaAverage 16 0x0000 Frequency Analyser average period of selected input pin. 0x144 FreqAnaCountInc 20 0x0000 0 Frequency Analyser counter increment amount. For each clock cycle no edge is detected on the selected input pin the accumulator is incremented by this amount. 0x148 FreqAnaCount 32 0x0000_0000 Frequency Analyser running counter (Working register) Miscellaneous 0x150 InterruptSrcSelect 10 0x3FF Interrupt source select. 1 bit per selected input. Determines whether the interrupt source is direct form the selected input pin or the deglitched version. Input pins are selected by the DeGlitchPinSelect register. 0 - Selected input direct 1 - Deglitched selected input 0x154 DebugSelect[8:2] 7 0x00 Debug address select. Indicates the address of the register to report on the gpio_cpu_data bus when it is not otherwise being used. 0x158-0x15C MotorMasterCount[1:0] 2 × 16 0x0000 Motor master clock counter values. Bus 0 - Master clock count 0 Bus 1 - Master clock count 1 Read Only registers 0x160 WakeUpInputMask 10 0x000 Indicates which deglitched inputs should be considered to generate the CPR wakeup. Active high 0x164 WakeUpLevel 1 0 Defines the level to detect on the masked GPIO inputs to generate a wakeup to the CPR 0 - Level 0 1 - Level 1 0x168 USBOverCurrentPinSelect 4 0x00 Selects which deglitched input should be used for the USB over current detect. 13.11.2.1 Supervisor and User Mode Access

The configuration registers block examines the CPU access type (cpu_acode signal) and determines if the access is allowed to that particular register, based on configured user access registers. If an access is not allowed the GPIO will issue a bus error by asserting the gpio_cpu_berr signal.

All supervisor and user program mode accesses will result in a bus error.

Access to the CpuIODirection, CpuIOOut and CpuIOIn is filtered by the CpuIOUserModeMask and CpuIOSuperModeMask registers. Each bit masks access to the corresponding bits in the CpuIO* registers for each mode, with CpuIOUserModeMask filtering user data mode access and CpuIOSuperModeMask filtering supervisor data mode access.

The addition of the CpuIOSuperModeMask register helps prevent potential conflicts between user and supervisor code read modify write operations. For example a conflict could exist if the user code is interrupted during a read modify write operation by a supervisor ISR which also modifies the CpuIO* registers.

An attempt to write to a disabled bit in user or supervisor mode will be ignored, and an attempt to read a disabled bit returns zero. If there are no user mode enabled bits then access is not allowed in user mode and a bus error will result. Similarly for supervisor mode.

When writing to the CpuIOOut register, the value being written is XORed with the current value in the CpuIOOut register, and the result is reflected on the GPIO pins.

The pseudocode for determining access to the CpuIOOut register is shown below. Similar code could be shown for the CpuIODirection and CpuIOIn registers. Note that when writing to CpuIODirection data is deposited directly and not XORed with the existing data (as in the CpuIOOut case).

if (cpu_acode == SUPERVISOR_DATA_MODE) then // supervisor mode if (CpuIOSuperModeMask[31:0] == 0 ) then // access is denied, and bus error gpio_cpu_berr = 1 elsif (cpu_rwn == 1) then // read mode (no filtering needed) gpio_cpu_data[31:0] = CpuIOOut[31:0] else // write mode, filtered by mask mask[31:0] = (cpu_dataout[31:0] & CpuIOSuperModeMask[31:0]) CpuIOOut[31:0] = (cpu_dataout[31:0] {circumflex over ( )} mask[31:0] ) //bitwise XOR operator elsif (cpu_acode == USER_DATA_MODE) then // user datamode if (CpuIOUserModeMask[31:0] == 0 ) then // access is denied, and bus error gpio_cpu_berr = 1 elsif (cpu_rwn == 1) then // read mode, filtered by mask gpio_cpu_data = ( CpuIOOut[31:0] & CpuIOUserModeMask[31:0]) else // write mode, filtered by mask mask[31:0] = (cpu_dataout[31:0] & CpuIOUserModeMask[31:0]) CpuIOOut[31:0] = (cpu_dataout[31:0] {circumflex over ( )} mask[31:0] ) //bitwise XOR operator else // access is denied, bus error gpio_cpu_berr = 1

Table 86 details the access modes allowed for registers in the GPIO block. In supervisor mode all registers are accessible. In user mode forbidden accesses will result in a bus error (gpio_cpu_berr asserted).

TABLE 86 GPIO supervisor and user access modes Register Address Registers Access Permitted 0x000-0x07C IOModeSelect[31:0] Supervisor data mode only 0x080-0x94 InputPinSelect[9:0] Supervisor data mode only CPU IO Control 0x0B0 CpuIOUserModeMask Supervisor data mode only 0x0B4 CpuIOSuperModeMask Supervisor data mode only 0x0B8 CpuIODirection CpuIOUserModeMask and CpuIOSuperModeMask filtered 0x0BC CpuIOOut CpuIOUserModeMask and CpuIOSuperModeMask filtered 0x0C0 CpuIOIn CpuIOUserModeMask and CpuIOSuperModeMask filtered 0x0C4 CpuDeGlitchUserModeMask Supervisor data mode only 0x0C8 CpuIOInDeglitch CpuDeGlitchUserModeMask filtered. Unrestricted Supervisor data mode access Deglitch control 0x0D0-0x0D4 DeGlitchCount[1:0] Supervisor data mode only 0x0D8-0x0DC DeGlitchClkSrc[1:0] Supervisor data mode only 0x0E0 DeGlitchSelect Supervisor data mode only Motor Control 0x0E4 MotorCtrlUserModeEnable Supervisor data mode only 0x0E8-0x0EC MotorMasterClkPeriod[1:0] MotorCtrlUserModeEnable enabled. 0x0F0 MotorMasterClkSrc MotorCtrlUserModeEnable enabled. 0x0F4-0x100 MotorCtrlConfig[3:0] MotorCtrlUserModeEnable enabled 0x104 MotorMasterClkSelect MotorCtrlUserModeEnable enabled 0x108 MotorMasterClockEnable MotorCtrlUserModeEnable enabled BLDC Motor Controllers 0x10C BLDCMode MotorCtrlUserModeEnable Enabled 0x110 BLDCDirection MotorCtrlUserModeEnable Enabled LED control 0x114 LEDCtrlUserModeEnable Supervisor data mode only 0x118-0x124 LEDDutySelect[3:0] LEDCtrlUserModeEnable[3:0] enabled Frequency Analyser 0x130 FreqAnaUserModeEnable Supervisor data mode only 0x134 FreqAnaPinSelect FreqAnaUserModeEnable enabled 0x138 FreqAnaPinFormSelect FreqAnaUserModeEnable enabled 0x13C FreqAnaLastPeriod FreqAnaUserModeEnable enabled 0x140 FreqAnaAverage FreqAnaUserModeEnable enabled 0x144 FreqAnaCountInc FreqAnaUserModeEnable enabled 0x148 FreqAnaCount FreqAnaUserModeEnable enabled Miscellaneous 0x150 InterruptSrcSelect Supervisor data mode only 0x154 DebugSelect[8:2] Supervisor data mode only 0x158-0x15C MotorMasterCount[1:0] Supervisor data mode only 0x160 WakeUpInputMask Supervisor data mode only 0x164 WakeUpLevel Supervisor data mode only 0x168 USBOverCurrentPinSelect Supervisor data mode only 13.11.3 GPIO Partition 13.11.4 IO Control

The IO control block connects the IO pin drivers to internal signalling based on configured setup registers and debug control signals.

// Output Control for (i=0; i<32 ; i++) { if (debug_cntrl[i] == 1) then // debug mode gpio_e[i] = 1;gpio_o[i] =debug_data_out[i] else // normal mode case io_mode_select[i] is 0 : gpio_e[i] =1 ;gpio_o[i] =led_ctrl[0]  // LED output 1 1 : gpio_e[i] =1 ;gpio_o[i] =led_ctrl[1]  // LED output 2 2 : gpio_e[i] =1 ;gpio_o[i] =led_ctrl[2]  // LED output 3 3 : gpio_e[i] =1 ;gpio_o[i] =led_ctrl[3]  // LED output 4 4 : gpio_e[i] =1 ;gpio_o[i] =motor_ctrl[0] // Stepper Motor Control 1 5 : gpio_e[i] =1 ;gpio_o[i] =motor_ctrl[1] // Stepper Motor Control 2 6 : gpio_e[i] =1 ;gpio_o[i] =motor_ctrl[2] // Stepper Motor Control 3 7 : gpio_e[i] =1 ;gpio_o[i] =motor_ctrl[3] // Stepper Motor Control 4 8 : gpio_e[i] =1 ;gpio_o[i] =bldc_ctrl[0][0]  // BLDC Motor Control 1,output 1 9 : gpio_e[i] =1 ;gpio_o[i] =bldc_ctrl[0][1]  // BLDC Motor Control 1,output 2 10: gpio_e[i] =1 ;gpio_o[i] =bldc_ctrl[0][2]  // BLDC Motor Control 1,output 3 11: gpio_e[i] =1 ;gpio_o[i] =bldc_ctrl[0][3]  // BLDC Motor Control 1,output 4 12: gpio_e[i] =1 ;gpio_o[i] =bldc_ctrl[0][4]  // BLDC Motor Control 1,output 5 13: gpio_e[i] =1 ;gpio_o[i] =bldc_ctrl[0][5]  // BLDC Motor Control 1,output 6 14: gpio_e[i] =1 ;gpio_o[i] =bldc_ctrl[1][0]  // BLDC Motor Control 2,output 1 15: gpio_e[i] =1 ;gpio_o[i] =bldc_ctrl[1][1]  // BLDC Motor Control 2,output 2 16: gpio_e[i] =1 ;gpio_o[i] =bldc_ctrl[1][2]  // BLDC Motor Control 2,output 3 17: gpio_e[i] =1 ;gpio_o[i] =bldc_ctrl[1][3]  // BLDC Motor Control 2,output 4 18: gpio_e[i] =1 ;gpio_o[i] =bldc_ctrl[1][4]  // BLDC Motor Control 2,output 5 19: gpio_e[i] =1 ;gpio_o[i] =bldc_ctrl[1][5]  // BLDC Motor Control 2,output 6 20: gpio_e[i] =1 ;gpio_o[i] =lss_gpio_clk[0] // LSS Clk 0 21: gpio_e[i] =1 ;gpio_o[i] =lss_gpio_clk[1] // LSS Clk 1 22: gpio_e[i] =lss_gpio_e[0] ;gpio_o[i] =lss_gpio_dout[0]; // LSS Data 0 gpio_lss_din[0] = gpio_i[i] 23: gpio_e[i] =lss_gpio_e[1] ;gpio_o[i] =lss_gpio_dout[1]; // LSS Data 1 gpio_lss_din[1] = gpio_i[i] 24: gpio_e[i] =isi_gpio_e[0] ;gpio_o[i] =isi_gpio_dout[0]; // ISI Control 1 gpio_isi_din[0] = gpio_i[i] 25: gpio_e[i] =isi_gpio_e[1] ;gpio_o[i] =isi_gpio_dout[1]; // ISI Control 2 gpio_isi_din[1] = gpio_i[i] 26: gpio_e[i] =isi_gpio_e[2] ;gpio_o[i] =isi_gpio_dout[2]; // ISI Control 3 gpio_isi_din[2] = gpio_i[i] 27: gpio_e[i] =isi_gpio_e[3] ;gpio_o[i] =isi_gpio_dout[3]; // ISI Control 4 gpio_isi_din[3] = gpio_i[i] 28: gpio_e[i] =cpu_io_dir[i] ;gpio_o[i] =cpu_io_out[i]; // CPU Direct 29: gpio e[i] =1 ;gpio o[i]  =usbh gpio power en // USB host power enable 30: gpio e[i] =0 ;gpio o[i] =0 // Input only mode end case // all gpio are always readable by the CPU cpu_io_in[i] = gpio_i[i]; }

The input selection pseudocode, for determining which pin connects to which de-glitch circuit.

for( i=0 ;i < 10 ; i++) { pin_num = input_pin_select[i] deglitch_input[i] = gpio_i[pin_num] }

-   -   The gpio_usbh_over_current output to the USB core is driven by a         selected deglitched input (configured by the         USBOverCurrentPinSelect register).     -   index =USBOverCurrentPinSelect         gpio_usbh_over_current=cpu_io_in_deglitch[index]         13.11.5 Wakeup Generator

The wakeup generator compares the deglitched inputs with the configured mask (WakeUpInputMask) and level (WakeUpLevel), and determines whether to generate a wakeup to the CPR block.

for (i =0;i<10; i++) { if (wakeup_level = 0) then // level 0 active wakeup = wakeup OR wakeup_input_mask[i] AND NOT cpu_io_in_deglitch[i] else // level 1 active wakeup = wakeup OR wakeup_input_mask[i] AND cpu_io_in_deglitch[i] } // assign the output gpio_cpr_wakeup = wakeup 13.11.6 LED Pulse Generator

The pulse generator logic consists of a 7-bit counter that is incremented on a 1 μs pulse from the timers block (tim_pulse[0]). The LED control signal is generated by comparing the count value with the configured duty cycle for the LED (led_duty_sel).

The logic is given by:

for (i=0 i<4 ;i++) { // for each LED pin // period divided into 8 segments period_div8 = cnt[6:4]; if (period_div8 < led_duty_sel[i]) then led_ctrl[i] = 1 else led_ctrl[i] = 0 } // update the counter every 1us pulse if (tim_pulse[0] == 1) then cnt ++ 13.11.7 Stepper Motor Control

The motor controller consists of 2 counters, and 4 phase generator logic blocks, one per motor control pin. The counters decrement each time a timing pulse (cnt_en) is received. The counters start at the configured clock period value (motor_mas_clk_period) and decrement to zero. If the counters are enabled (via motor_mas_clk_enable), the counters will automatically restart at the configured clock period value, otherwise they will wait until the counters are re-enabled.

-   -   The timing pulse period is one of pclk, 1 μs, 100 μs, 1 ms         depending on the motor_mas_clk_sel signal. The counters are used         to derive the phase and duty cycle of each motor control pin.

// decrement logic if (cnt_en == 1) then if ((mas_cnt == 0) AND (motor_mas_clk_enable == 1)) then mas_cnt = motor_mas_clk_period[15:0] elsif ((mas_cnt == 0) AND (motor_mas_clk_enable == 0)) then mas_cnt = 0 else mas_cnt −− else // hold the value mas_cnt = mas_cnt

The phase generator block generates the motor control logic based on the selected clock generator (motor_mas_clk_sel) the motor control high transition point (curr_motor_ctrl_high) and the motor control low transition point (curr_motor_ctrl_low).

The phase generator maintains current copies of the motor_ctrl_config configuration value (motor_ctrl_config[31:16] becomes curr_motor_ctrl_high and motor_ctrl_config[15:0] becomes curr_motor_ctrl_low). It updates these values to the current register values when it is safe to do so without causing a glitch on the output motor pin.

Note that when reprogramming the motor_ctrl_config register to reorder the sequence of the transition points (e.g changing from low point less than high point to low point greater than high point and vice versa) care must taken to avoid introducing glitching on the output pin.

There are 4 instances one per motor control pin.

The logic is given by:

// select the input counter to use if (motor_mas_clk_sel == 1) then count = mas_cnt[1] else count = mas_cnt[0] // Generate the phase and duty cycle if (count == curr_motor_ctrl_low) then motor_ctrl = 0 elsif (count == curr_motor_ctrl_high) then motor_ctrl = 1 else motor_ctrl = motor_ctrl // remain the same // update the current registers at period boundary if (count == 0) then curr_motor_ctrl_high = motor_ctrl_config[31:16] // update to new high value curr_motor_ctrl_low = motor_ctrl_config[15:0] // update to new high value 13.11.8 Input Deglitch

The input deglitch logic rejects input states of duration less than the configured number of time units (deglitch_cnt), input states of greater duration are reflected on the output cpu_io_in_deglitch. The time units used (either pclk, 1 μs, 100 μs, 1 ms) by the deglitch circuit is selected by the deglitch_clk_src bus.

There are 2 possible sets of deglitch_cnt and deglitch_clk_src that can be used to deglitch the input pins. The values used are selected by the deglitch_sel signal.

There are 10 deglitch circuits in the GPIO. Any GPIO pin can be connected to a deglitch circuit. Pins are selected for deglitching by the InputPinSelect registers.

Each selected input can be used to generate an interrupt. The interrupt can be generated from the raw input signal (deglitch_input) or a deglitched version of the input (cpu_io_in_deglitch). The interrupt source is selected by the interrupt_src_select signal.

The counter logic is given by

if (deglitch_input != deglitch_input_delay) then cnt = deglitch_cnt output_en = 0 elsif (cnt == 0 ) then cnt = cnt output_en = 1 elsif (cnt_en == 1) then cnt −− output_en = 0 13.11.9 Frequency Analyser

The frequency analyser block monitors a selected deglitched input (cpu_io_in_deglitch) or a direct selected input (deglitch_input) and detects positive edges. The selected input is configured by FreqAnaPinSelect and FreqAnaPinFormSel registers. Between successive positive edges detected on the input it increments a counter (FreqAnaCount) by a programmed amount (FreqAnaCountInc) on each clock cycle. When a positive edge is detected the FreqAnaLastPeriod register is updated with the top 16 bits of the counter and the counter is reset. The frequency analyser also maintains a running average of the FreqAnaLastPeriod register. Each time a positive edge is detected on the input the FreqAnaAverage register is updated with the new calculated FreqAnaLastPeriod. The average is calculated as ⅞ the current value plus ⅛ of the new value. The FreqAnaLastPeriod, FreqAnaCount and FreqAnaAverage registers can be written to by the CPU.

The pseudocode is given by

if ((pin == 1) AND pin_delay ==0 )) then // positive edge detected freq_ana_lastperiod[15:0] = freq_ana_count[31:16] freq_ana_average[15:0] = freq_ana_average[15:0] − freq_ana_average [15:3] + freq_ana_lastperiod[15:3] freq_ana_count[15:0] = 0 else freq_ana_count[31:0] = freq_ana_count[31:0]  + freq_ana_count_inc[19:0] // implement the configuration register write if (wr_last_en == 1) then freq_ana_lastperiod = wr_data elsif (wr_average_en == 1 ) then freq_ana_average = wr_data elsif (wr_freq_count_en == 1) then freq_ana_count = wr_data 13.11.10 BLDC Motor Controller

The BLDC controller logic is identical for both instances, only the input connections are different. The logic implements the truth table shown in Table. The six q outputs are combinationally based on the direction, ha, hb, hc and pwm inputs. The direction input has 2 possible sources selected by the mode, the pseudocode is as follows

// determine if in internal or external direction mode if (mode == 1) then // internal mode direction = int_direction else // external mode direction = ext_direction 14 Interrupt Controller Unit (ICU)

The interrupt controller accepts up to N input interrupt sources, determines their priority, arbitrates based on the highest priority and generates an interrupt request to the CPU. The ICU complies with the interrupt acknowledge protocol of the CPU. Once the CPU accepts an interrupt (i.e. processing of its service routine begins) the interrupt controller will assert the next arbitrated interrupt if one is pending.

Each interrupt source has a fixed vector number N, and an associated configuration register, IntReg[N]. The format of the IntReg[N] register is shown in Table 87 below.

TABLE 87 IntReg[N] register format Field bit(s) Description Priority 3:0 Interrupt priority Type 5:4 Determines the triggering conditions for the interrupt 00 - Positive edge 10 - Negative edge 01 - Positive level 11 - Negative level Mask 6 Mask bit. 1 - Interrupts from this source are enabled, 0 - Interrupts from this source are disabled. Note that there may be additional masks in operation at the source of the interrupt. Reserved 31:7 Reserved. Write as 0.

Once an interrupt is received the interrupt controller determines the priority and maps the programmed priority to the appropriate CPU priority levels, and then issues an interrupt to the CPU. The programmed interrupt priority maps directly to the LEON CPU interrupt levels. Level 0 is no interrupt. Level 15 is the highest interrupt level.

14.1 Interrupt Preemption

With standard LEON pre-emption an interrupt can only be pre-empted by an interrupt with a higher priority level. If an interrupt with the same priority level (1 to 14) as the interrupt being serviced becomes pending then it is not acknowledged until the current service routine has completed. Note that the level 15 interrupt is a special case, in that the LEON processor will continue to take level 15 interrupts (i.e re-enter the ISR) as long as level 15 is asserted on the icu_cpu_ilevel. Level 0 is also a special case, in that LEON consider level 0 interrupts as no interrupt, and will not issue an acknowledge when level 0 is presented on the icu_cpu_ilevel bus.

Thus when pre-emption is required, interrupts should be programmed to different levels as interrupt priorities of the same level have no guaranteed servicing order. Should several interrupt sources be programmed with the same priority level, the lowest value interrupt source will be serviced first and so on in increasing order.

The interrupt is directly acknowledged by the CPU and the ICU automatically clears the pending bit of the lowest value pending interrupt source mapped to the acknowledged interrupt level.

All interrupt controller registers are only accessible in supervisor data mode. If the user code wishes to mask an interrupt it must request this from the supervisor and the supervisor software will resolve user access levels.

14.2 Interrupt Sources

The mapping of interrupt sources to interrupt vectors (and therefore IntReg[N] registers) is shown in Table 88 below. Please refer to the appropriate section of this specification for more details of the interrupt sources.

TABLE 88 Interrupt sources vector table Vector Source Description 0 Timers WatchDog Timer Update request 1 Timers Generic Timer 1 interrupt 2 Timers Generic Timer 2 interrupt 3 PCU PEP Sub-system Interrupt- TE finished band 4 PCU PEP Sub-system Interrupt- LBD finished band 5 PCU PEP Sub-system Interrupt- CDU finished band 6 PCU PEP Sub-system Interrupt- CDU error 7 PCU PEP Sub-system Interrupt- PCU finished band 8 PCU PEP Sub-system Interrupt- PCU Invalid address interrupt 9 PHI PEP Sub-system Interrupt- PHI Line Sync Interrupt 10 PHI PEP Sub-system Interrupt- PHI Buffer underrun 11 PHI PEP Sub-system Interrupt- PHI Page finished 12 PHI PEP Sub-system Interrupt- PHI Print ready 13 SCB USB Host interrupt 14 SCB USB Device interrupt 15 SCB ISI interrupt 16 SCB DMA interrupt 17 LSS LSS interrupt, LSS interface 0 interrupt request 18 LSS LSS interrupt, LSS interface 1 interrupt request 19-28 GPIO GPIO general purpose interrupts 29 Timers Generic Timer 3 interrupt 14.3 Implentation 14.3.1 Definitions of I/O

TABLE 89 Interrupt Controller Unit I/O definition Port name Pins I/O Description Clocks and Resets Pclk 1 In System Clock prst_n 1 In System reset, synchronous active low CPU interface cpu_adr[7:2] 6 In CPU address bus. Only 6 bits are required to decode the address space for the ICU block cpu_dataout[31:0] 32 In Shared write data bus from the CPU icu_cpu_data[31:0] 32 Out Read data bus to the CPU cpu_rwn 1 In Common read/not-write signal from the CPU cpu_icu_sel 1 In Block select from the CPU. When cpu_icu_sel is high both cpu_adr and cpu_dataout are valid icu_cpu_rdy 1 Out Ready signal to the CPU. When icu_cpu_rdy is high it indicates the last cycle of the access. For a write cycle this means cpu_dataout has been registered by the ICU block and for a read cycle this means the data on icu_cpu_data is valid. icu_cpu_ilevel[3:0] 4 Out Indicates the priority level of the current active interrupt. cpu_iack 1 In Interrupt request acknowledge from the LEON core. cpu_icu_ilevel[3:0] 4 In Interrupt acknowledged level from the LEON core icu_cpu_berr 1 Out Bus error signal to the CPU indicating an invalid access. cpu_acode[1:0] 2 In CPU Access Code signals. These decode as follows: 00 - User program access 01 - User data access 10 - Supervisor program access 11 - Supervisor data access icu_cpu_debug_valid 1 Out Debug Data valid on icu_cpu_data bus. Active high Interrupts tim_icu_wd_irq 1 In Watchdog timer interrupt signal from the Timers block tim_icu_irq[2:0] 3 In Generic timer interrupt signals from the Timers block gpio_icu_irq[9:0] 10 In GPIO pin interrupts usb_icu_irq[1:0] 2 In USB host and device interrupts from the SCB Bit 0 - USB Host interrupt Bit 1 - USB Device interrupt isi_icu_irq 1 In ISI interrupt from the SCB dma_icu_irq 1 In DMA interrupt from the SCB lss_icu_irq[1:0] 2 In LSS interface interrupt request cdu_finishedband 1 In Finished band interrupt request from the CDU cdu_icu_jpegerror 1 In JPEG error interrupt from the CDU lbd_finishedband 1 In Finished band interrupt request from the LBD te_finishedband 1 In Finished band interrupt request from the TE pcu_finishedband 1 In Finished band interrupt request from the PCU pcu_icu_address_invalid 1 In Invalid address interrupt request from the PCU phi_icu_underrun 1 In Buffer underrun interrupt request from the PHI phi_icu_page_finish 1 In Page finished interrupt request from the PHI phi_icu_print_rdy 1 In Print ready interrupt request from the PHI phi_icu_linesync_int 1 In Line sync interrupt request from the PHI 14.3.2 Configuration Registers

The configuration registers in the ICU are programmed via the CPU interface. Refer to section 11.4 on page 69 for a description of the protocol and timing diagrams for reading and writing registers in the ICU. Note that since addresses in SoPEC are byte aligned and the CPU only supports 32-bit register reads and writes, the lower 2 bits of the CPU address bus are not required to decode the address space for the ICU. When reading a register that is less than 32 bits wide zeros should be returned on the upper unused bit(s) of icu_pcu_data. Table 90 lists the configuration registers in the ICU block.

The ICU block will only allow supervisor data mode accesses (i.e. cpu_acode[1:0]=SUPERVISOR_DATA). All other accesses will result in icu_cpu_berr being asserted.

TABLE 90 ICU Register Map Address ICU_base+ Register #bits Reset Description 0x00-0x74 IntReg[29:0] 30 × 7 0x00 Interrupt vector configuration register 0x88 IntClear 30 0x0000_0000 Interrupt pending clear register. If written with a one it clears corresponding interrupt Bits[30:0] - Interrupts sources 30 to 0 (Reads as zero) 0x90 IntPending 30 0x0000_0000 Interrupt pending register. (Read Only) Bits[30:0]- Interrupts sources 30 to 0 0xA0 IntSource 5 0x1F Indicates the interrupt source of the last acknowl- edged interrupt. The NoInterrupt value is defined as all bits set to one. (Read Only) 0xC0 DebugSelect[7:2] 6 0x00 Debug address select. Indicates the address of the register to report on the icu_cpu_data bus when it is not otherwise being used. 14.3.3 ICU Partition 114.3.4 Interrupt Detect

The ICU contains multiple instances of the interrupt detect block, one per interrupt source. The interrupt detect block examines the interrupt source signal, and determines whether it should generate request pending (int_pend) based on the configured interrupt type and the interrupt source conditions. If the interrupt is not masked the interrupt will be reflected to the interrupt arbiter via the int_active signal. Once an interrupt is pending it remains pending until the interrupt is accepted by the CPU or it is level sensitive and gets removed. Masking a pending interrupt has the effect of removing the interrupt from arbitration but the interrupt will still remain pending.

When the CPU accepts the interrupt (using the normal ISR mechanism), the interrupt controller automatically generates an interrupt clear for that interrupt source (cpu_int_clear). Alternatively if the interrupt is masked, the CPU can determine pending interrupts by polling the IntPending registers. Any active pending interrupts can be cleared by the CPU without using an ISR via the IntClear registers.

Should an interrupt clear signal (either from the interrupt clear unit or the CPU) and a new interrupt condition happen at the same time, the interrupt will remain pending. In the particular case of a level sensitive interrupt, if the level remains the interrupt will stay active regardless of the clear signal.

The logic is shown below:

mask = int_config[6] type = int_config[5:4] int_pend  = last_int_pend // the last pending interrupt // update the pending FF // test for interrupt condition if (type == NEG_LEVEL) then int_pend = NOT(int_src) elsif (type == POS_LEVEL) int_pend = int_src elsif ((type == POS_EDGE ) AND (int_src == 1) AND (last_int_src == 0)) int_pend = 1 elsif ((type == NEG_EDGE ) AND (int_src == 0) AND (last_int_src == 1)) int_pend = 1 elsif ((int_clear == 1 )OR (cpu_int_clear==1)) then int_pend = 0 else int_pend = last_int_pend // stay the same as before // mask the pending bit if (mask == 1) then int_active = int_pend else int_active = 0 // assign the registers last_int_src = int_src last_int_pend = int_pend 14.3.5 Interrupt Arbiter

The interrupt arbiter logic arbitrates a winning interrupt request from multiple pending requests based on configured priority. It generates the interrupt to the CPU by setting icu_cpu_ilevel to a non-zero value. The priority of the interrupt is reflected in the value assigned to icu_cpu_ilevel, the higher the value the higher the priority, 15 being the highest, and 0 considered no interrupt.

// arbitrate with the current winner int_ilevel = 0 for (i=0;i<30;i++) { if ( int_active[i] == 1) then { if (int_config[i][3:0] > win_int_ilevel[3:0] ) then win_int_ilevel[3:0] = int_config[i][3:0] } } } // assign the CPU interrupt level int_ilevel = win_int_ilevel[3:0] 14.3.6 Interrupt Clear Unit

The interrupt clear unit is responsible for accepting an interrupt acknowledge from the CPU, determining which interrupt source generated the interrupt, clearing the pending bit for that source and updating the IntSource register.

When an interrupt acknowledge is received from the CPU, the interrupt clear unit searches through each interrupt source looking for interrupt sources that match the acknowledged interrupt level (cpu_icu_ilevel) and determines the winning interrupt (lower interrupt source numbers have higher priority). When found the interrupt source pending bit is cleared and the IntSource register is updated with the interrupt source number.

The LEON interrupt acknowledge mechanism automatically disables all other interrupts temporarily until it has correctly saved state and jumped to the ISR routine. It is the responsibility of the ISR to re-enable the interrupts. To prevent the IntSource register indicating the incorrect source for an interrupt level, the ISR must read and store the IntSource value before re-enabling the interrupts via the Enable Traps (ET) field in the Processor State Register (PSR) of the LEON.

See section 11.9 on page 104 for a complete description of the interrupt handling procedure. After reset the state machine remains in Idle state until an interrupt acknowledge is received from the CPU (indicated by cpu_iack). When the acknowledge is received the state machine transitions to the Compare state, resetting the source counter (cnt) to the number of interrupt sources.

While in the Compare state the state machine cycles through each possible interrupt source in decrementing order. For each active interrupt source the programmed priority (int_priority[cnt][3:0]) is compared with the acknowledged interrupt level from the CPU (cpu_icu_ilevel), if they match then the interrupt is considered the new winner. This implies the last interrupt source checked has the highest priority, e.g interrupt source zero has the highest priority and the first source checked has the lowest priority. After all interrupt sources are checked the state machine transitions to the IntClear state, and updates the int_source register on the transition.

Should there be no active interrupts for the acknowledged level (e.g. a level sensitive interrupt was removed), the IntSource register will be set to NoInterrupt. NoInterrupt is defined as the highest possible value that IntSource can be set to (in this case 0x1F), and the state machine will return to Idle.

The exact number of compares performed per clock cycle is dependent the number of interrupts, and logic area to logic speed trade-off, and is left to the implementer to determine. A comparison of all interrupt sources must complete within 8 clock cycles (determined by the CPU acknowledge hardware).

When in the IntClear state the state machine has determined the interrupt source to clear (indicated by the int_source register). It resets the pending bit for that interrupt source, transitions back to the Idle state and waits for the next acknowledge from the CPU.

The minimum time between successive interrupt acknowledges from the CPU is 8 cycles.

15 Timers Block (TIM)

The Timers block contains general purpose timers, a watchdog timer and timing pulse generator for use in other sections of SoPEC.

15.1 Watchdog Timer

The watchdog timer is a 32 bit counter value which counts down each time a timing pulse is received. The period of the timing pulse is selected by the WatchDogUnitSel register. The value at any time can be read from the WatchDogTimer register and the counter can be reset by writing a non-zero value to the register. When the counter transitions from 1 to 0, a system wide reset will be triggered as if the reset came from a hardware pin.

The watchdog timer can be polled by the CPU and reset each time it gets close to 1, or alternatively a threshold (WatchDogIntThres) can be set to trigger an interrupt for the watchdog timer to be serviced by the CPU. If the WatchDogIntThres is set to N, then the interrupt will be triggered on the N to N−1 transition of the WatchDogTimer. This interrupt can be effectively masked by setting the threshold to zero. The watchdog timer can be disabled, without causing a reset, by writing zero to the WatchDogTimer register.

15.2 Timing Pulse Generator

The timing block contains a timing pulse generator clocked by the system clock, used to generate timing pulses of programmable periods. The period is programmed by accessing the TimerStartValue registers. Each pulse is of one system clock duration and is active high, with the pulse period accurate to the system clock frequency. The periods after reset are set to 1 us, 100 us and 100 ms.

The timing pulse generator also contains a 64-bit free running counter that can be read or reset by accessing the FreeRunCount registers. The free running counter can be used to determine elapsed time between events at system clock accuracy or could be used as an input source in low-security random number generator.

15.3 Generic Timers

SoPEC contains 3 programmable generic timing counters, for use by the CPU to time the system. The timers are programmed to a particular value and count down each time a timing pulse is received. When a particular timer decrements from 1 to 0, an interrupt is generated. The counter can be programmed to automatically restart the count, or wait until re-programmed by the CPU. At any time the status of the counter can be read from GenCntValue, or can be reset by writing to GenCntValue register. The auto-restart is activated by setting the GenCntAuto register, when activated the counter restarts at GenCntStartValue. A counter can be stopped or started at any time, without affecting the contents of the GenCntValue register, by writing a 1 or 0 to the relevent GenCntEnable register.

15.4 Implementation

15.4.1 Definitions of I/O

TABLE 91 Timers block I/O definition Port name Pins I/O Description Clocks and Resets Pclk 1 In System Clock prst_n 1 In System reset, synchronous active low tim_pulse[2:0] 3 Out Timers block generated timing pulses, each one pclk wide 0 - Nominal 1 μs pulse 1 - Nominal 100 μs pulse 2 - Nominal 10 ms pulse CPU interface cpu_adr[6:2] 5 In CPU address bus. Only 5 bits are required to decode the address space for the ICU block cpu_dataout[31:0] 32 In Shared write data bus from the CPU tim_cpu_data[31:0] 32 Out Read data bus to the CPU cpu_rwn 1 In Common read/not-write signal from the CPU cpu_tim_sel 1 In Block select from the CPU. When cpu_tim_sel is high both cpu_adr and cpu_dataout are valid tim_cpu_rdy 1 Out Ready signal to the CPU. When tim_cpu_rdy is high it indicates the last cycle of the access. For a write cycle this means cpu_dataout has been registered by the TIM block and for a read cycle this means the data on tim_cpu_data is valid. tim_cpu_berr 1 Out Bus error signal to the CPU indicating an invalid access. cpu_acode[1:0] 2 In CPU Access Code signals. These decode as follows: 00 - User program access 01 - User data access 10 - Supervisor program access 11 - Supervisor data access tim_cpu_debug_valid 1 Out Debug Data valid on tim_cpu_data bus. Active high Miscellaneous tim_icu_wd_irq 1 Out Watchdog timer interrupt signal to the ICU block tim_icu_irq[2:0] 3 Out Generic timer interrupt signals to the ICU block tim_cpr_reset_n 1 Out Watch dog timer system reset. 15.4.2 Timers Sub-Block Partition 15.4.3 Watchdog Timer

The watchdog timer counts down from pre-programmed value, and generates a system wide reset when equal to one. When the counter passes a pre-programmed threshold (wdog_tim_thres) value an interrupt is generated (tim_icu_wd_irq) requesting the CPU to update the counter. Setting the counter to zero disables the watchdog reset. In supervisor mode the watchdog counter can be written to or read from at any time, in user mode access is denied. Any accesses in user mode will generate a bus error.

The counter logic is given by

if (wdog_wen == 1) then wdog_tim_cnt = write_data // load new data elsif ( wdog_tim_cnt == 0) then wdog_tim_cnt = wdog_tim_cnt // count disabled elsif ( cnt_en == 1 ) then wdog_tim_cnt−− else wdog_tim_cnt = wdog_tim_cnt

The timer decode logic is

if (( wdog_tim_cnt == wdog_tim_thres) AND (wdog_tim_cnt != 0 )AND (cnt_en == 1)) then tim_icu_wd_irq = 1 else tim_icu_wd_irq = 0 // reset generator logic if (wdog_tim_cnt == 1) AND (cnt_en == 1) then tim_cpr_reset_n = 0 else tim_cpr_reset_n = 1 15.4.4 Generic Timers

The generic timers block consists of 3 identical counters. A timer is set to a pre-configured value (GenCntStartValue) and counts down once per selected timing pulse (gen_unit_sel). The timer can be enabled or disabled at any time (gen_tim_en), when disabled the counter is stopped but not cleared. The timer can be set to automatically restart (gen_tim_auto) after it generates an interrupt. In supervisor mode a timer can be written to or read from at any time, in user mode access is determined by the GenCntUserModeEnable register settings.

The counter logic is given by

if (gen_wen == 1) then gen_tim_cnt = write_data elsif (( cnt_en == 1 )AND (gen_tim_en == 1 )) then if ( gen_tim_cnt == 1) OR ( gen_tim_cnt == 0) then // counter may need re-starting if (gen_tim_auto == 1) then gen_tim_cnt = gen_tim_cnt_st_value else gen_tim_cnt = 0 // hold count at zero else gen_tim_cnt−− else gen_tim_cnt = gen_tim_cnt

The decode logic is

if (gen_tim_cnt == 1)AND ( cnt_en == 1 )AND (gen_tim_en == 1 ) then tim_icu_irq = 1 else time_icu_irq = 0 15.4.5 Timing Pulse Generator

The timing pulse generator contains a general free running 64-bit timer and 3 timing pulse generators producing timing pulses of one cycle duration with a programmable period. The period is programmed by changed the TimerStartValue registers, but have a nominal starting period of 1 μs, 100 μs and 1 ms. In supervisor mode the free running timer register can be written to or read from at any time, in user mode access is denied. The status of each of the timers can be read by accessing the PulseTimerStatus registers in supervisor mode. Any accesses in user mode will result in a bus error.

15.4.5.1 Free Run Timer

The increment logic block increments the timer count on each clock cycle. The counter wraps around to zero and continues incrementing if overflow occurs. When the timing register (FreeRunCount) is written to, the configuration registers block will set the free_run_wen high for a clock cycle and the value on write_data will become the new count value. If free_run_wen[1] is 1 the higher 32 bits of the counter will be written to, otherwise if free_run_wen[0] the lower 32 bits are written to. It is the responsibility of software to handle these writes in a sensible manner.

The increment logic is given by

if (free_run_wen[1] == 1) then free_run_cnt[63:32] = write_data elsif (free_run_wen[0] == 1) then free_run_cnt[31:0] = write_data else free_run_cnt ++ 15.4.5.2 Pulse Timers

The pulse timer logic generates timing pulses of 1 clock cycle length and programmable period. Nominally they generate pulse periods of 1 μs, 100 μs and 1 ms. The logic for timer 0 is given by:

// Nominal 1us generator if (pulse_0_cnt == 0 ) then pulse_0_cnt = timer_start_value[0] tim_pulse[0] = 1 else pulse_0_cnt −− tim_pulse[0] = 0

The logic for timer 1 is given by:

// 100us generator if ((pulse_1_cnt == 0) AND (tim_pulse[0] == 1)) then pulse_1_cnt = timer_start_value[1] tim_pulse[1] = 1 elsif (tim_pulse[0] == 1) then pulse_1_cnt −− tim_pulse[1] = 0 else pulse_1_cnt = pulse_1_cnt tim_pulse[1] = 0

The logic for the timer 2 is given by:

// 10ms generator if ((pulse_2_cnt == 0 ) AND (tim_pulse[1] == 1)) then pulse_2_cnt = timer_start_value[2] tim_pulse[2] = 1 elsif (tim_pulse[1] == 1) then pulse_2_cnt −− tim_pulse[2] = 0 else pulse_2_cnt = pulse_2_cnt tim_pulse[2] = 0 15.4.6 Configuration Registers

The configuration registers in the TIM are programmed via the CPU interface. Refer to section 11.4.3 on page 69 for a description of the protocol and timing diagrams for reading and writing registers in the TIM. Note that since addresses in SoPEC are byte aligned and the CPU only supports 32-bit register reads and writes, the lower 2 bits of the CPU address bus are not required to decode the address space for the TIM. When reading a register that is less than 32 bits wide zeros should be returned on the upper unused bit(s) of tim_pcu_data. Table 92 lists the configuration registers in the TIM block.

TABLE 92 Timers Register Map Address TIM_base+ Register #bits Reset Description 0x00 WatchDogUnitSel 2 0x0 Specifies the units used for the watchdog timer: 0 - Nominal 1 μs pulse 1 - Nominal 100 μs pulse 2 - Nominal 10 ms pulse 3 - pclk 0x04 WatchDogTimer 32 0xFFFF_FFFF Specifies the number of units to count before watchdog timer triggers. 0x08 WatchDogIntThres 32 0x0000_0000 Specifies the threshold value below which the watchdog timer issues an interrupt 0x0C-0x10 FreeRunCount[1:0]  2 × 32 0x0000_0000 Direct access to the free running counter register. Bus 0 - Access to bits 31-0 Bus 1 - Access to bits 63-32 0x14 to 0x1C GenCntStartValue[2:0]  3 × 32 0x0000_0000 Generic timer counter start value, number of units to count before event 0x20 to 0x28 GenCntValue[2:0]  3 × 32 0x0000_0000 Direct access to generic timer counter registers 0x2C to 0x34 GenCntUnitSel[2:0] 3 × 2 0x0 Generic counter unit select. Selects the timing units used with corresponding counter: 0 - Nominal1 μs pulse 1 - Nominal100 μs pulse 2 - Nominal 10 ms pulse 3 - pclk 0x38 to 0x40 GenCntAuto[2:0] 3 × 1 0x0 Generic counter auto re-start select. When high timer automatically restarts, otherwise timer stops. 0x44 to 0x4C GenCntEnable[2:0] 3 × 1 0x0 Generic counter enable. 0 - Counter disabled 1 - Counter enabled 0x50 GenCntUserModeEnable 3 0x0 User Mode Access enable to generic timer configuration register. When 1 user access is enabled. Bit 0 - Generic timer 0 Bit 1 - Generic timer 1 Bit 2 - Generic timer 2 0x54 to 0x5C TimerStartValue[2:0] 3 × 8 0x7F, Timing pulse generator start value. 0x63, Indicates the start value for each 0x63 timing pulse timers. For timer 0 the start value specifies the timer period in pclk cycles − 1. For timer 1 the start value specifies the timer period in timer 0 intervals − 1. For timer 2 the start value specifies the timer period in timer 1 intervals − 1. Nominally the timers generate pulses at 1 us, 100 us and 10 ms intervals respectively. 0x60 DebugSelect[6:2] 5 0x00 Debug address select. Indicates the address of the register to report on the tim_cpu_data bus when it is not otherwise being used. Read Only Registers 0x64 PulseTimerStatus 24 0x00 Current pulse timer values, and pulses 7:0 - Timer 0 count 15:8 - Timer 1 count 23:16 - Timer 2 count 24 - Timer 0 pulse 25 - Timer 1 pulse 26 - Timer 2 pulse 15.4.6.1 Supervisor and User Mode Access

The configuration registers block examines the CPU access type (cpu_acode signal) and determines if the access is allowed to that particular register, based on configured user access registers. If an access is not allowed the block will issue a bus error by asserting the tim_cpu_berr signal.

The timers block is fully accessible in supervisor data mode, all registers can written to and read from. In user mode access is denied to all registers in the block except for the generic timer configuration registers that are granted user data access. User data access for a generic timer is granted by setting corresponding bit in the GenCntUserModeEnable register. This can only be changed in supervisor data mode. If a particular timer is granted user data access then all registers for configuring that timer will be accessible. For example if timer 0 is granted user data access the GenCntStartValue[0], GenCntUnitSel[0], GenCntAuto[0], GenCntEnable[0] and GenCntValue[0] registers can all be written to and read from without any restriction.

Attempts to access a user data mode disabled timer configuration register will result in a bus error. Table 93 details the access modes allowed for registers in the TIM block. In supervisor data mode all registers are accessable. All forbidden accesses will result in a bus error (tim_cpu_berr asserted).

TABLE 93 TIM supervisor and user access modes Register Address Registers Access Permission 0x00 WatchDogUnitSel Supervisor data mode only 0x04 WatchDogTimer Supervisor data mode only 0x08 WatchDogIntThres Supervisor data mode only 0x0C-0x10 FreeRunCount Supervisor data mode only 0x14 GenCntStartValue[0] GenCntUserModeEnable[0] 0x18 GenCntStartValue[1] GenCntUserModeEnable[1] 0x1C GenCntStartValue[2] GenCntUserModeEnable[2] 0x20 GenCntValue[0] GenCntUserModeEnable[0] 0x24 GenCntValue[1] GenCntUserModeEnable[1] 0x28 GenCntValue[2] GenCntUserModeEnable[2] 0x2C GenCntUnitSel[0] GenCntUserModeEnable[0] 0x30 GenCntUnitSel[1] GenCntUserModeEnable[1] 0x34 GenCntUnitSel[2] GenCntUserModeEnable[2] 0x38 GenCntAuto[0] GenCntUserModeEnable[0] 0x3C GenCntAuto[1] GenCntUserModeEnable[1] 0x40 GenCntAuto[2] GenCntUserModeEnable[2] 0x44 GenCntEnable[0] GenCntUserModeEnable[0] 0x48 GenCntEnable[1] GenCntUserModeEnable[1] 0x4C GenCntEnable[2] GenCntUserModeEnable[2] 0x50 GenCntUserModeEnable Supervisor data mode only 0x54-0x5C TimerStartValue[2:0] Supervisor data mode only 0x60 DebugSelect Supervisor data mode only 0x64 PulseTimerStatus Supervisor data mode only 16 Clocking, Power and Reset (CPR)

The CPR block provides all of the clock, power enable and reset signals to the SoPEC device.

16.1 Powerdown Modes

The CPR block is capable of powering down certain sections of the SoPEC device. When a section is powered down (i.e. put in sleep mode) no state is retained (except the PSS storage), the CPU must re-initialize the section before it can be used again.

For the purpose of powerdown the SoPEC device is divided into sections:

TABLE 94 Powerdown sectioning Section Block Print Engine Pipeline PCU Subsystem (Section 0) CDU CFU LBD SFU TE TFU HCU DNC DWU LLU PHI CPU-DRAM (Section 1) DRAM CPU/MMU DIU TIM ROM LSS PSS ICU ISI Subsystem (Section 2) ISI (SCB) DMA Ctrl (SCB) GPIO USB Subsystem (Section 3) USB (SCB)

Note that the CPR block is not located in any section. All configuration registers in the CPR block are clocked by an ungateable clock and have special reset conditions.

16.1.1 Sleep Mode

Each section can be put into sleep mode by setting the corresponding bit in the SleepModeEnable register. To re-enable the section the sleep mode bit needs to be cleared and then the section should be reset by writing to the relevant bit in the ResetSection register. Each block within the section should then be re-configured by the CPU.

If the CPU system (section 1) is put into sleep mode, the SoPEC device will remain in sleep mode until a system level reset is initiated from the reset pin, or a wakeup reset by the SCB block as a result of activity on either the USB or ISI bus. The watchdog timer cannot reset the device as it is in section 1 also, and will be in sleep mode.

If the CPU and ISI subsystem are in sleep mode only a reset from the USB or a hardware reset will re-activate the SoPEC device.

If all sections are put into sleep mode, then only a system level reset initiated by the reset pin will re-activate the SoPEC device.

Like all software resets in SoPEC the ResetSection register is active-low i e. a 0 should be written to each bit position requiring a reset. The ResetSection register is self-reseting.

16.1.2 Sleep Mode Powerdown Procedure

When powering down a section, the section may retain it's current state (although not gauranteed to). It is possible when powering back up a section that inconsistencies between interface state machines could cause incorrect operation. In order to prevent such condition from happening, all blocks in a section must be disabled before powering down. This will ensure that blocks are restored in a benign state when powered back up.

In the case of PEP section units setting the Go bit to zero will disable the block. The DRAM subsystem can be effectively disabled by setting the RotationSync bit to zero, and the SCB system disabled by setting the DMAAccessEn bits to zero turning off the DMA access to DRAM. Other CPU subsystem blocks without any DRAM access do not need to be disabled.

16.2 Reset Source

The SoPEC device can be reset by a number of sources. When a reset from an internal source is initiated the reset source register (ResetSrc) stores the reset source value. This register can then be used by the CPU to determine the type of boot sequence required.

16.3 Clock Relationship

The crystal oscillator excites a 32 MHz crystal through the xtalin and xtalout pins. The 32 MHz output is used by the PLL to derive the master VCO frequency of 960 MHz. The master clock is then divided to produce 320 MHz clock (clk320), 160 MHz clock (clk160) and 48 MHz (clk48) clock sources.

The phase relationship of each clock from the PLL will be defined. The relationship of internal clocks clk320, clk48 and clk160 to xtalin will be undefined.

At the output of the clock block, the skew between each pclk domain (pclk_section[2:0] and jclk) should be within skew tolerances of their respective domains (defined as less than the hold time of a D-type flip flop).

The skew between doclk and pclk should also be less than the skew tolerances of their respective domains.

The usbclk is derived from the PLL output and has no relationship with the other clocks in the system and is considered asynchronous.

16.4 PLL Control

The PLL in SoPEC can be adjusted by programming the PLLRangeA, PLLRangeB, PLLTunebits and PLLMult registers. If these registers are changed by the CPU the values are not updated until the PLLUpdate register is written to. Writing to the PLLUpdate register triggers the PLL control state machine to update the PLL configuration in a safe way. When an update is active (as indicated by PLLUpdate register) the CPU must not change any of the configuration registers, doing so could cause the PLL to lose lock indefintely, requiring a hardware reset to recover. Configuring the PLL registers in an inconsistent way can also cause the PLL to lose lock, care must taken to keep the PLL configuration within specified parameters.

The VCO frequency of the PLL is calculated by the number of divider in the feedback path. PLL output A is used as the feedback source. VCOfreq=REFCLK×PLLMult×PLLRangeA×External divider VCOfreq=32×3×10×1=960 Mhz.

In the default PLL setup, PLLMult is set to 3, PLLRangeA is set to 3 which corresponds to a divide by 10, PLLRangeB is set to 5 which corresponds to a divide by 3. PLLouta=VCOfreq/PLLRangeA=960 Mhz/10=96 Mhz PLLoutb=VCOfreq/PLLRangeB=960 Mhz/3=320 Mhz

See [16] for complete PLL setup parameters.

16.5 Implementation

16.5.1 Definitions of I/O

TABLE 95 CPR I/O definition Port name Pins I/O Description Clocks and Resets Xtalin 1 In Crystal input, di- rect from IO pin. Xtalout 1 Inout Crystal output, direct to IO pin. pclk_section[3:0] 4 Out System clocks for each section Doclk 1 Out Data out clock (2× pclk) for the PHI block Jclk 1 Out Gated version of system clock used to clock the JPEG decoder core in the CDU Usbclk 1 Out USB clock, nominally at 48 Mhz jclk_enable 1 In Gating signal for jclk. When 1 jclk is enabled reset_n 1 In Reset signal from the reset_n pin usb_cpr_reset_n 1 In Reset signal from the USB block isi_cpr_reset_n 1 In Reset signal from the ISI block tim_cpr_reset_n 1 In Reset signal from watch dog timer. gpio_cpr_wakeup 1 In SoPEC wake up from the GPIO, active high. prst_n_section[3:0] 4 Out System resets for each section, synchronous active low dorst_n 1 Out Reset for PHI block, synchronous to doclk jrst_n 1 Out Reset for JPEG decoder core in CDU block, synchronous to jclk usbrst_n 1 Out Reset for the USB block, synchronous to usbclk CPU interface cpu_adr[5:2] 3 In CPU address bus. Only 4 bits are required to decode the address space for the CPR block cpu_dataout[31:0] 32 In Shared write data bus from the CPU cpr_cpu_data[31:0] 32 Out Read data bus to the CPU cpu_rwn 1 In Common read/not- write signal from the CPU cpu_cpr_sel 1 In Block select from the CPU. When cpu_cpr_sel is high both cpu_adr and cpu_dataout are valid cpr_cpu_rdy 1 Out Ready signal to the CPU. When cpr_cpu_rdy is high it indicates the last cycle of the access. For a write cycle this means cpu_dataout has been registered by the block and for a read cycle this means the data on cpr_cpu_data is valid. cpr_cpu_berr 1 Out Bus error signal to the CPU indicating an invalid access. cpu_acode[1:0] 2 In CPU Access Code signals. These decode as follows: 00 - User program access 01 - User data access 10 - Supervisor program access 11 - Supervisor data access cpr_cpu_debug_valid 1 Out Debug Data valid on cpr_cpu_data bus. Active high 16.5.2 Configuration Registers

The configuration registers in the CPR are programmed via the CPU interface. Refer to section 11.4 on page 69 for a description of the protocol and timing diagrams for reading and writing registers in the CPR. Note that since addresses in SoPEC are byte aligned and the CPU only supports 32-bit register reads and writes, the lower 2 bits of the CPU address bus are not required to decode the address space for the CPR. When reading a register that is less than 32 bits wide zeros should be returned on the upper unused bit(s) of cpr_pcu_data. Table 96 lists the configuration registers in the CPR block.

-   -   The CPR block will only allow supervisor data mode accesses         (i.e. cpu_acode[1:0]=SUPERVISOR_DATA). All other accesses will         result in cpr_cpu_berr being asserted.

TABLE 96 CPR Register Map Address CPR_base+ Register #bits Reset Description 0x00 SleepModeEnable 4 0x0^(a) Sleep Mode enable, when high a section of logic is put into powerdown. Bit 0 - Controls section 0 Bit 1 - Controls section 1 Bit 2 - Controls section 2 Bit 3 - Controls section 3 Note that the SleepModeEnable register has special reset conditions. See Section 16.5.6 for details 0x04 ResetSrc 5 0x1^(a) Reset Source register, indicating the source of the last reset (or wake-up) Bit 0 - External Reset Bit 1 - USB wakeup reset Bit 2 - ISI wakeup reset Bit 3 - Watchdog timer reset Bit 4 - GPIO wake-up (Read Only Register) 0x08 ResetSection 4 0xF Active-low synchronous reset for each section, self-resetting. Bit 0 - Controls section 0 Bit 1 - Controls section 1 Bit 2 - Controls section 2 Bit 3 - Controls section 3 0x0C DebugSelect[5:2] 4 0x0 Debug address select. Indicates the address of the register to report on the cpr_cpu_data bus when it is not otherwise being used. PLL Control 0x10 PLLTuneBits 10 0x3BC PLL tuning bits 0x14 PLLRangeA 4 0x3 PLLOUT A frequency selector (defaults to 60 Mhz to 125 Mhz) 0x18 PLLRangeB 3 0x5 PLLOUT B frequency selector (defaults to 200 Mhz to 400 Mhz) 0x1C PLLMultiplier 5 0x03 PLL multiplier selector, defaults to refclk × 3 0x20 PLLUpdate 1 0x0 PLL update control. A write (of any value) to this register will cause the PLL to lose lock for ~100 us. Reading the register indicates the status of the update. 0 - PLL update complete 1 - PLL update active No writes to PLLTuneBits, PLLRangeA, PLL- RangeB, PLLMultiplier or PLLUpdate are allowed while the PLL update is active. ^(a)Reset value depends on reset source. External reset shown. 16.5.3 CPR Sub-Block Partition 16.5.4 Reset_n deglitch

The external reset_n signal is deglitched for about 1 μs. reset_n must maintain a state for 1 us second before the state is passed into the rest of the device. All deglitch logic is clocked on bufrefclk.

16.5.5 Sync Reset

The reset synchronizer retimes an asynchronous reset signal to the clock domain that it resets. The circuit prevents the inactive edge of reset occurring when the clock is rising

16.5.6 Reset Generator Logic

The reset generator logic is used to determine which clock domains should be reset, based on configured reset values (reset_section_n), the external reset (reset_n), watchdog timer reset (tim_cpr_reset_n), the USB reset (usb_cpr_reset_n), the GPIO wakeup control (gpio_cpr_wakeup) and the ISI reset (isi_cpr_reset_n). The reset direct from the IO pin (reset_n) is synchronized and de-glitched before feeding the reset logic.

All resets are lengthened to at least 16 pclk cycles, regardless of the duration of the input reset. The clock for a particular section must be running for the reset to have an effect. The clocks to each section can be enabled/disabled using the SleepModeEnable register.

Resets from the ISI or USB block reset everything except its own section (section 2 or 3).

TABLE 97 Reset domains Reset signal Domain reset_dom[0] Section 0 pclk domain (PEP) reset_dom[1] Section 1 pclk domain (CPU) reset_dom[2] Section 2 pclk domain (ISI) reset_dom[3] Section 3 usbclk/pclk domain (USB) reset_dom[4] doclk domain reset_dom[5] jclk domain

The logic is given by

if (reset_dg_n == 0) then reset_dom[5:0] = 0x00 // reset everything reset_src[4:0] = 0x01 cfg_reset_n = 0 sleep_mode_en[3:0] = 0x0 // re-awaken all sections elsif (tim_cpr_reset_n == 0) then reset_dom[5:0]  = 0x00  // reset everything except CPR config reset_src[4:0] = 0x08 cfg_reset_n = 1 // CPR config stays the same sleep_mode_en[1]  = 0  // re-awaken section 1 only (awake already) elsif (usb_cpr_reset_n == 0) then reset_dom[5:0]  = 0x08   // all except USB domain + CPR config reset_src[4:0] = 0x02 cfg_reset_n = 1 // CPR config stays the same sleep_mode_en[1]  = 0  // re-awaken section 1 only, section 3 is awake elsif (isi_cpr_reset_n == 0) then reset_dom[5:0]  = 0x04   // all except ISI domain + CPR config reset_src[4:0] = 0x04 cfg_reset_n = 1 // CPR config stays the same sleep_mode_en[1]  = 0  // re-awaken section 1 only, section 2 is awake elsif (gpio_cpr_wakeup = 1) then reset_dom[5:0] = 0x3C // PEP and CPU sections only reset_src[4:0] = 0x10 cfg_reset_n = 1 // CPR config stays the same sleep_mode_en[1]  = 0  // re-awaken section 1 only, section 2 is awake else // propagate resets from reset section register reset_dom[5:0] = 0x3F // default to on cfg_reset_n = 1 // CPR cfg registers are not in any section sleep_mode_en[3:0]  = sleep_mode_en[3:0] // stay the same by default if (reset_section_n[0] == 0) then reset_dom[5] = 0   // jclk domain reset_dom[4] = 0   // doclk domain reset_dom[0] = 0   // pclk section 0 domain if (reset_section_n[1] == 0) then reset_dom[1] = 0   // pclk section 1 domain if (reset_section_n[2] == 0) then reset_dom[2] = 0 // pclk section 2 domain (ISI) if (reset_section_n[3] == 0) then reset_dom[3] = 0   // USB domain 16.5.7 Sleep Logic

The sleep logic is used to generate gating signals for each of SoPECs clock domains. The gate enable (gate_dom) is generated based on the configured sleep_mode_en and the internally generated jclk_enable signal.

The logic is given by

// clock gating for sleep modes gate_dom[5:0] = 0x0  // default to all clocks on if (sleep_mode_en[0] == 1) then  // section 0 sleep gate_dom[0] = 1 // pclk section 0 gate_dom[4] = 1 // doclk domain gate_dom[5] = 1 // jclk domain if (sleep_mode_en[1] == 1) then  // section 1 sleep gate_dom[1] = 1 // pclk section 1 if (sleep_mode_en[2] == 1) then  // section 2 sleep gate_dom[2] = 1 // pclk section 2 if (sleep_mode_en[3] == 1) then  // section 3 sleep gate_dom[3] = 1 // usb section 3 // the jclk can be turned off by CDU signal if (jclk_enable == 0) then gate_dom[5] = 1

The clock gating and sleep logic is clocked with the master_pclk clock which is not gated by this logic, but is synchronous to other pclk_section and jclk domains.

Once a section is in sleep mode it cannot generate a reset to restart the device. For example if section 1 is in sleep mode then the watchdog timer is effectively disabled and cannot trigger a reset.

16.5.8 Clock Gate Logic

The clock gate logic is used to safely gate clocks without generating any glitches on the gated clock. When the enable is high the clock is active otherwise the clock is gated

16.5.9 Clock Generator Logic

The clock generator block contains the PLL, crystal oscillator, clock dividers and associated control logic. The PLL VCO frequency is at 960 MHz locked to a 32 MHz refclk generated by the crystal oscillator. In test mode the xtalin signal can be driven directly by the test clock generator, the test clock will be reflected on the refclk signal to the PLL.

16.5.9.1 Clock Divider A

The clock divider A block generates the 48 MHz clock from the input 96 MHz clock (pllouta) generated by the PLL. The divider is enabled only when the PLL has acquired lock.

16.5.9.2 Clock Divider B

The clock divider B block generates the 160 MHz clocks from the input 320 MHz clock (plloutb) generated by the PLL. The divider is enabled only when the PLL has acquired lock.

16.5.9.3 PLL Control State Machine

The PLL will go out of lock whenever pll_reset goes high (the PLL reset is the only active high reset in the device) or if the configuration bits pll_rangea, pll_rangeb, pll_mult, pll_tune are changed. The PLL control state machine ensures that the rest of the device is protected from glitching clocks while the PLL is being reset or it's configuration is being changed.

In the case of a hardware reset (the reset is deglitched), the state machine first disables the output clocks (via the clk_gate signal), it then holds the PLL in reset while its configuration bits are reset to default values. The state machine then releases the PLL reset and waits approx. 100 us to allow the PLL to regain lock. Once the lock time has elapsed the state machine re-enables the output clocks and resets the remainder of the device via the reset_dg_n signal.

When the CPU changes any of the configuration registers it must write to the PLLupdate register to allow the state machine to update the PLL to the new configuration setup. If a PLLUpdate is detected the state machine first gates the output clocks. It then holds the PLL in reset while the PLL configuration registers are updated. Once updated the PLL reset is released and the state machine waits approx 100 us for the PLL to regain lock before re-enabling the output clocks. Any write to the PLLUpdate register will cause the state machine to perform the update operation regardless of whether the configuration values changed or not.

All logic in the clock generator is clocked on bufrefclk which is always an active clock regardless of the state of the PLL.

17 ROM Block

17.1 Overview

The ROM block interfaces to the CPU bus and contains the SoPEC boot code. The ROM block consists of the CPU bus interface, the ROM macro and the ChipID macro. The current ROM size is 16 KBytes implemented as a 4096×32 macro. Access to the ROM is not cached because the CPU enjoys fast (no more than one cycle slower than a cache access), unarbitrated access to the ROM. Each SoPEC device is required to have a unique ChipID which is set by blowing fuses at manufacture. IBM's 300 mm ECID macro and a custom 112-bit ECID macro are used to implement the ChipID offering 224-bits of laser fuses. The exact number of fuse bits to be used for the ChipID will be determined later but all bits are made available to the CPU. The ECID macros allows all 224 bits to be read out in parallel and the ROM block will make all 224 bits available in the FuseChipID[N] registers which are readable by the CPU in supervisor mode only.

17.2 Boot Operation

The are two boot scenarIOs for the SoPEC device namely after power-on and after being awoken from sleep mode. When the device is in sleep mode it is hoped that power will actually be removed from the DRAM, CPU and most other peripherals and so the program code will need to be freshly downloaded each time the device wakes up from sleep mode. In order to reduce the wakeup boot time (and hence the perceived print latency) certain data items are stored in the PSS block (see section 18). These data items include the SHA-1 hash digest expected for the program(s) to be downloaded, the master/slave SoPEC id and some configuration parameters. All of these data items are stored in the PSS by the CPU prior to entering sleep mode. The SHA-1 value stored in the PSS is calculated by the CPU by decrypting the signature of the downloaded program using the appropriate public key stored in ROM. This compute intensive decryption only needs to take place once as part of the power-on boot sequence—subsequent wakeup boot sequences will simply use the resulting SHA-1 digest stored in the PSS. Note that the digest only needs to be stored in the PSS before entering sleep mode and the PSS can be used for temporary storage of any data at all other times.

The CPU is expected to be in supervisor mode for the entire boot sequence described by the pseudocode below. Note that the boot sequence has not been finalised but is expected to be close to the following:

if (ResetSrc == 1) then // Reset was a power-on reset configure_sopec // need to configure peris (USB, ISI, DMA, ICU etc.) // Otherwise reset was a wakeup reset so peris etc. were already configured PAUSE: wait until IrqSemaphore != 0 // i.e. wait until an interrupt has been serviced if (IrqSemaphore == DMAChan0Msg) then parse_msg(DMAChan0MsgPtr) // this routine will parse the message and take any // necessary action e.g. programming the DMAChannel1 registers elsif (IrqSemaphore == DMAChan1Msg) then // program has been downloaded CalculatedHash = gen_sha1(ProgramLocn, ProgramSize) if (ResetSrc == 1) then ExpectedHash = sig_decrypt(ProgramSig,public_key) else ExpectedHash = PSSHash if (ExpectedHash == CalculatedHash) then jmp(PrgramLocn) // transfer control to the downloaded program else send_host_msg(“Program Authentication Failed”) goto PAUSE: elsif (IrqSemaphore == timeout) then // nothing has happened if (ResetSrc == 1) then sleep_mode( ) // put SoPEC into sleep mode to be woken up by USB/ISI activity else // we were woken up but nothing happened reset_sopec(PowerOnReset) else goto PAUSE

The boot code places no restrictions on the activity of any programs downloaded and authenticated by it other than those imposed by the configuration of the MMU i.e. the principal function of the boot code is to authenticate that any programs downloaded by it are from a trusted source. It is the responsibility of the downloaded program to ensure that any code it downloads is also authenticated and that the system remains secure. The downloaded program code is also responsible for setting the SoPEC ISIId (see section 12.5 for a description of the ISIID) in a multi-SoPEC system. See the “SoPEC Security Overview” document [9] for more details of the SoPEC security features.

17.3 Implementation

17.3.1 Definitions of I/O

TABLE 98 ROM Block I/O Port name Pins I/O Description Clocks and Resets prst_n 1 In Global reset. Synchronous to pclk, active low. Pclk 1 In Global clock CPU Interface cpu_adr[14:2] 13 In CPU address bus. Only 13 bits are required to decode the address space for this block. rom_cpu_data[31:0] 32 Out Read data bus to the CPU cpu_rwn 1 In Common read/not-write signal from the CPU cpu_acode[1:0] 2 In CPU Access Code signals. These decode as follows: 00 - User program access 01 - User data access 10 - Supervisor program access 11 - Supervisor data access cpu_rom_sel 1 In Block select from the CPU. When cpu_rom_sel is high cpu_adr is valid rom_cpu_rdy 1 Out Ready signal to the CPU. When rom_cpu_rdy is high it indicates the last cycle of the access. For a read cycle this means the data on rom_cpu_data is valid. rom_cpu_berr 1 Out ROM bus error signal to the CPU indicating an invalid access. 17.3.2 Configuration Registers

The ROM block will only allow read accesses to the FuseChipID registers and the ROM with supervisor data space permissions (i.e. cpu_acode[1:0]=11). Write accesses with supervisor data space permissions

will have no effect. All other accesses with will result in rom_cpu_berr being asserted. The CPU subsystem bus slave interface is described in more detail in section 9.4.3.

TABLE 99 ROM Block Register Map Address ROM_base+ Register #bits Reset Description 0x4000 FuseChipID0 32 n/a Value of corresponding fuse bits 31 to 0 of the IBM 112-bit ECID macro. (Read only) 0x4004 FuseChipID1 32 n/a Value of corresponding fuse bits 63 to 32 of the IBM 112-bit ECID macro. (Read only) 0x4008 FuseChipID2 32 n/a Value of corresponding fuse bits 95 to 64 of the IBM 112-bit ECID macro. (Read only) 0x400C FuseChipID3 16 n/a Value of corresponding fuse bits 111 to 96 of the IBM 112-bit ECID macro. (Read only) 0x4010 FuseChipID4 32 n/a Value of corresponding fuse bits 31 to 0 of the Custom 112-bit ECID macro. (Read only) 0x4014 FuseChipID5 32 n/a Value of corresponding fuse bits 63 to 32 of the Custom 112-bit ECID macro. (Read only) 0x4018 FuseChipID6 32 n/a Value of corresponding fuse bits 95 to 64 of the Custom 112-bit ECID macro. (Read only) 0x401C FuseChipID7 16 n/a Value of corresponding fuse bits 111 to 96 of the Custom 112-bit ECID macro. (Read only) 17.3.3 Sub-Block Partition

IBM offer two variants of their ROM macros; A high performance version (ROMHD) and a low power version (ROMLD). It is likely that the low power version will be used unless some implementation issue requires the high performance version. Both versions offer the same bit density. The sub-block partition diagram below does not include the clocking and test signals for the ROM or ECID macros. The CPU subsystem bus interface is described in more detail in section 11.4.3.

TABLE 100 ROM Block internal signals Port name Width Description Clocks and Resets prst_n 1 Global reset. Synchronous to pclk, active low. Pclk 1 Global clock Internal Signals rom_adr[11:0] 12 ROM address bus rom_sel 1 Select signal to the ROM macro instructing it to access the location at rom_adr rom_oe 1 Output enable signal to the ROM block rom_data[31:0] 32 Data bus from the ROM macro to the CPU bus interface rom_dvalid 1 Signal from the ROM macro indicating that the data on rom_data is valid for the address on rom_adr fuse_data[31:0] 32 Data from the FuseChipID[N] register addressed by fuse_reg_adr fuse_reg_adr[2:0] 3 Indicates which of the FuseChipID registers is being addressed Sub-Block Signal Definition 18 Power Safe Storage (PSS) Block 18.1 Overview

The PSS block provides 128 bytes of storage space that will maintain its state when the rest of the SoPEC device is in sleep mode. The PSS is expected to be used primarily for the storage of decrypted signatures associated with downloaded programmed code but it can also be used to store any information that needs to survive sleep mode (e.g. configuration details). Note that the signature digest only needs to be stored in the PSS before entering sleep mode and the PSS can be used for temporary storage of any data at all other times.

Prior to entering sleep mode the CPU should store all of the information it will need on exiting sleep mode in the PSS. On emerging from sleep mode the boot code in ROM will read the ResetSrc register in the CPR block to determine which reset source caused the wakeup. The reset source information indicates whether or not the PSS contains valid stored data, and the PSS data determines the type of boot sequence to execute. If for any reason a full power-on boot sequence should be performed (e.g. the printer driver has been updated) then this is simply achieved by initiating a full software reset.

Note that a reset or a powerdown (powerdown is implemented by clock gating) of the PSS block will not clear the contents of the 128 bytes of storage. If clearing of the PSS storage is required, then the CPU must write to each location individually.

18.2 Implementation

The storage area of the PSS block will be implemented as a 128-byte register array. The array is located from PSS_base through to PSS_base+0x7F in the address map. The PSS block will only allow read or write accesses with supervisor data space permissions (i.e. cpu_acode[1:0]=11). All other accesses will result in pss_cpu_berr being asserted. The CPU subsystem bus slave interface is described in more detail in section 11.4.3.

18.2.1 Definitions of I/O

TABLE 101 PSS Block I/O Port name Pins I/O Description Clocks and Resets prst_n 1 In Global reset. Synchronous to pclk, active low. Pclk 1 In Global clock CPU Interface cpu_adr[6:2] 5 In CPU address bus. Only 5 bits are required to decode the address space for this block. cpu_dataout[31:0] 32 In Shared write data bus from the CPU pss_cpu_data[31:0] 32 Out Read data bus to the CPU cpus_rwn 1 In Common read/not-write signal from the CPU cpu_acode[1:0] 2 In CPU Access Code signals. These decode as follows: 00 - User program access 01 - User data access 10 - Supervisor program access 11 - Supervisor data access cpu_pss_sel 1 In Block select from the CPU. When cpu_pss_sel is high both cpu_adr and cpu_dataout are valid pss_cpu_rdy 1 Out Ready signal to the CPU. When pss_cpu_rdy is high it indi- cates the last cycle of the access. For a read cycle this means the data on pss_cpu_data is valid. pss_cpu_berr 1 Out PSS bus error signal to the CPU indicating an invalid access. 19 Low Speed Serial Interface (LSS) 19.1 Overview

The Low Speed Serial Interface (LSS) provides a mechanism for the internal SoPEC CPU to communicate with external QA chips via two independent LSS buses. The LSS communicates through the GPIO block to the QA chips. This allows the QA chip pins to be reused in multi-SoPEC environments. The LSS Master system-level interface is illustrated in FIG. 75. Note that multiple QA chips are allowed on each LSS bus.

19.2 QA Communication

The SoPEC data interface to the QA Chips is a low speed, 2 pin, synchronous serial bus. Data is transferred to the QA chips via the lss_data pin synchronously with the lss_clk pin. When the lss_clk is high the data on lss_data is deemed to be valid. Only the LSS master in SoPEC can drive the lss_clk pin, this pin is an input only to the QA chips. The LSS block must be able to interface with an open-collector pull-up bus. This means that when the LSS block should transmit a logical zero it will drive 0 on the bus, but when it should transmit a logical 1 it will leave high-impedance on the bus (i.e. it doesn't drive the bus). If all the agents on the LSS bus adhere to this protocol then there will be no issues with bus contention.

The LSS block controls all communication to and from the QA chips. The LSS block is the bus master in all cases. The LSS block interprets a command register set by the SoPEC CPU, initiates transactions to the QA chip in question and optionally accepts return data. Any return information is presented through the configuration registers to the SoPEC CPU. The LSS block indicates to the CPU the completion of a command or the occurrence of an error via an interrupt. The LSS protocol can be used to communicate with other LSS slave devices (other than QA chips). However should a LSS slave device hold the clock low (for whatever reason), it will be in violation of the LSS protocol and is not supported. The LSS clock is only ever driven by the LSS master.

19.2.1 Start and Stop Conditions

All transmissions on the LSS bus are initiated by the LSS master issuing a START condition and terminated by the LSS master issuing a STOP condition. START and STOP conditions are always generated by the LSS master. As illustrated in FIG. 76, a START condition corresponds to a high to low transition on lss_data while lss_clk is high. A STOP condition corresponds to a low to high transition on lss_data while lss_clk is high.

19.2.2 Data Transfer

Data is transferred on the LSS bus via a byte orientated protocol. Bytes are transmitted serially. Each byte is sent most significant bit (MSB) first through to least significant bit (LSB) last. One clock pulse is generated for each data bit transferred. Each byte must be followed by an acknowledge bit.

The data on the lss_data must be stable during the HIGH period of the lss_clk clock. Data may only change when lss_clk is low. A transmitter outputs data after the falling edge of lss_clk and a receiver inputs the data at the rising edge of lss_clk. This data is only considered as a valid data bit at the next lss_clk falling edge provided a START or STOP is not detected in the period before the next lss_clk falling edge. All clock pulses are generated by the LSS block. The transmitter releases the lss_data line (high) during the acknowledge clock pulse (ninth clock pulse). The receiver must pull down the lss_data line during the acknowledge clock pulse so that it remains stable low during the HIGH period of this clock pulse.

Data transfers follow the format shown in FIG. 77. The first byte sent by the LSS master after a START condition is a primary id byte, where bits 7-2 form a 6-bit primary id (0 is a global id and will address all QA Chips on a particular LSS bus), bit 1 is an even parity bit for the primary ID, and bit 0 forms the read/write sense. Bit 0 is high if the following command is a read to the primary id given or low for a write command to that id. An acknowledge is generated by the QA chip(s) corresponding to the given id (if such a chip exists) by driving the lss_data line low synchronous with the LSS master generated ninth lss_clk.

19.2.3 Write Procedure

The protocol for a write access to a QA Chip over the LSS bus is illustrated in FIG. 79 below. The LSS master in SoPEC initiates the transaction by generating a START condition on the LSS bus. It then transmits the primary id byte with a 0 in bit 0 to indicate that the following command is a write to the primary id. An acknowledge is generated by the QA chip corresponding to the given primary id. The LSS master will clock out M data bytes with the slave QA Chip acknowledging each successful byte written. Once the slave QA chip has acknowledged the M^(th) data byte the LSS master issues a STOP condition to complete the transfer. The QA chip gathers the M data bytes together and interprets them as a command. See QA Chip Interface Specification for more details on the format of the commands used to communicate with the QA chip[8]. Note that the QA chip is free to not acknowledge any byte transmitted. The LSS master should respond by issuing an interrupt to the CPU to indicate this error. The CPU should then generate a STOP condition on the LSS bus to gracefully complete the transaction on the LSS bus.

19.2.4 Read Procedure

The LSS master in SoPEC initiates the transaction by generating a START condition on the LSS bus. It then transmits the primary id byte with a 1 in bit 0 to indicate that the following command is a read to the primary id. An acknowledge is generated by the QA chip corresponding to the given primary id. The LSS master releases the lss_data bus and proceeds to clock the expected number of bytes from the QA chip with the LSS master acknowledging each successful byte read. The last expected byte is not acknowledged by the LSS master. It then completes the transaction by generating a STOP condition on the LSS bus. See QA Chip Interface Specification for more details on the format of the commands used to communicate with the QA chip[8].

19.3 Implementation

A block diagram of the LSS master is given in FIG. 80. It consists of a block of configuration registers that are programmed by the CPU and two identical LSS master units that generate the signalling protocols on the two LSS buses as well as interrupts to the CPU. The CPU initiates and terminates transactions on the LSS buses by writing an appropriate command to the command register, writes bytes to be transmitted to a buffer and reads bytes received from a buffer, and checks the sources of interrupts by reading status registers.

19.3.1 Definitions of IO

TABLE 102 LSS IO pins definitions Port name Pins I/O Description Clocks and Resets Pclk 1 In System Clock prst_n 1 In System reset, synchronous active low CPU Interface cpu_rwn 1 In Common read/not-write signal from the CPU cpu_adr[6:2] 5 In CPU address bus. Only 5 bits are required to decode the address space for this block cpu_dataout[31:0] 32 In Shared write data bus from the CPU cpu_acode[1:0] 2 In CPU access code signals. cpu_acode[0] - Program (0)/ Data (1) access cpu_acode[1] - User (0)/ Supervisor (1) access cpu_lss_sel 1 In Block select from the CPU. When cpu_lss_sel is high both cpu_adr and cpu_dataout are valid lss_cpu_rdy 1 Out Ready signal to the CPU. When lss_cpu_rdy is high it indicates the last cycle of the access. For a write cycle this means cpu_dataout has been registered by the LSS block and for a read cycle this means the data on lss_cpu_data is valid. lss_cpu_berr 1 Out LSS bus error signal to the CPU. lss_cpu_data[31:0] 32 Out Read data bus to the CPU lss_cpu_debug_valid 1 Out Active high. Indicates the presence of valid debug data on lss_cpu_data. GPIO for LSS buses lss_gpio_dout[1:0] 2 Out LSS bus data output Bit 0 - LSS bus 0 Bit 1 - LSS bus 1 gpio_lss_din[1:0] 2 In LSS bus data input Bit 0 - LSS bus 0 Bit 1 - LSS bus 1 lss_gpio_e[1:0] 2 Out LSS bus data output enable, active high Bit 0 - LSS bus 0 Bit 1 - LSS bus 1 lss_gpio_clk[1:0] 2 Out LSS bus clock output Bit 0 - LSS bus 0 Bit 1 - LSS bus 1 ICU interface lss_icu_irq[1:0] 2 Out LSS interrupt requests Bit 0 - interrupt associated with LSS bus 0 Bit 1 - interrupt associated with LSS bus 1 19.3.2 Configuration Registers

The configuration registers in the LSS block are programmed via the CPU interface. Refer to section 11.4 on page 69 for the description of the protocol and timing diagrams for reading and writing registers in the LSS block. Note that since addresses in SoPEC are byte aligned and the CPU only supports 32-bit register reads and writes, the lower 2 bits of the CPU address bus are not required to decode the address space for the LSS block. Table 103 lists the configuration registers in the LSS block. When reading a register that is less than 32 bits wide zeros should be returned on the upper unused bit(s) of lss_cpu_data.

The input cpu_acode signal indicates whether the current CPU access is supervisor, user, program or data. The configuration registers in the LSS block can only be read or written by a supervisor data access, i.e. when cpu_acode equals b11. If the current access is a supervisor data access then the LSS responds by asserting lss_cpu_rdy for a single clock cycle.

If the current access is anything other than a supervisor data access, then the LSS generates a bus error by asserting lss_cpu_berr for a single clock cycle instead of lss_cpu_rdy as shown in section 11.4 on page 69. A write access will be ignored, and a read access will return zero.

TABLE 103 LSS Control Registers Address (LSS_base+) Register #bits Reset Description Control registers 0x00 Reset 1 0x1 A write to this register causes a reset of the LSS. 0x04 LssClockHighLowDuration 16 0x00C8 Lss_clk has a 50:50 duty cycle, this register defines the period of lss_clk by means of specifying the duration (in pclk cycles) that lss_clk is low (or high). The reset value specifies transmission over the LSS bus at a nominal rate of 400 kHz, corresponding to a low (or high) duration of 200 pclk (160 Mhz) cycles. Register should not be set to values less than 8. 0x08 LssClocktoDataHold 6 0x3 Specifies the number of pclk cycles that Data must remain valid for after the falling edge of lss_clk. Minimum value is 3 cycles, and must to programmed to be less than LssClockHighLowDuration. LSS bus 0 registers 0x10 Lss0IntStatus 3 0x0 LSS bus 0 interrupt status registers Bit 0 - command completed successfully Bit 1 - error during processing of command, not -acknowledge received after transmission of primary id byte on LSS bus 0 Bit 2 - error during processing of command, not -acknowledge received after transmission of data byte on LSS bus 0 All the bits in Lss0IntStatus are cleared when the Lss0Cmd register gets written to. (Read only register) 0x14 Lss0CurrentState 4 0x0 Gives the current state of the LSS bus 0 state machine. (Read only register). (Encoding will be specified upon state machine implementation) 0x18 Lss0Cmd 21 0x00_0000 Command register defining sequence of events to perform on LSS bus 0 before interrupting CPU. A write to this register causes all the bits in the Lss0IntStatus register to be cleared as well as generating a lss0_new_cmd pulse. 0x1C-0x2C Lss0Buffer[4:0] 5 × 32 0x0000_0000 LSS Data buffer. Should be filled with transmit data before transmit command, or read data bytes received after a valid read command. LSS bus 1 registers 0x30 Lss1IntStatus 3 0x0 LSS bus 1 interrupt status registers Bit 0 - command completed successfully Bit 1 - error during processing of command, not -acknowledge received after transmission of primary id byte on LSS bus 1 Bit 2 - error during processing of command, not -acknowledge received after transmission of data byte on LSS bus 1 All the bits in Lss1IntStatus are cleared when the Lss1Cmd register gets written to. (Read only register) 0x34 Lss1CurrentState 4 0x0 Gives the current state of the LSS bus 1 state machine. (Read only register) (Encoding will be specified upon state machine implementation) 0x38 Lss1Cmd 21 0x00_0000 Command register defining sequence of events to perform on LSS bus 1 before interrupting CPU. A write to this register causes all the bits in the Lss1IntStatus register to be cleared as well as generating a lss1_new_cmd pulse. 0x3C-0x4C Lss1Buffer[4:0] 5 × 32 0x0000_0000 LSS Data buffer. Should be filled with transmit data before transmit command, or read data bytes received after a valid read command. Debug registers 0x50 LssDebugSel[6:2] 5 0x00 Selects register for debug output. This value is used as the input to the register decode logic instead of cpu_adr[6:2] when the LSS block is not being accessed by the CPU, i.e. when cpu_lss_sel is 0. The output lss_cpu_debug_valid is asserted to indicate that the data on lss_cpu_data is valid debug data. This data can be multiplexed onto chip pins during debug mode. 19.3.2.1 LSS Command Registers

The LSS command registers define a sequence of events to perform on the respective LSS bus before issuing an interrupt to the CPU. There is a separate command register and interrupt for each LSS bus. The format of the command is given in Table 104. The CPU writes to the command register to initiate a sequence of events on an LSS bus. Once the sequence of events has completed or an error has occurred, an interrupt is sent back to the CPU.

Some example commands are:

-   -   a single START condition (Start=1, IdByteEnable=0, RdWrEnable=0,         Stop=0)     -   a single STOP condition (Start=0, IdByteEnable=0, RdWrEnable=0,         Stop=1)     -   a START condition followed by transmission of the id byte         (Start=1, IdByteEnable=1, RdWrEnable=0, Stop=0, IdByte contains         primary id byte)     -   a write transfer of 20 bytes from the data buffer (Start=0,         IdByteEnable=0, RdWrEnable=1, RdWrSense=0, Stop=0,         TxRxByteCount=20)     -   a read transfer of 8 bytes into the data buffer (Start=0,         IdByteEnable=0, RdWrEnable=1, RdWrSense=1, ReadNack=0, Stop=0,         TxRxByteCount=8)     -   a complete read transaction of 16 bytes (Start=1,         IdByteEnable=1, RdWrEnable=1, RdWrSense=1, ReadNack=1, Stop=1,         IdByte contains primary id byte, TxRxByteCount=16), etc.

The CPU can thus program the number of bytes to be transmitted or received (up to a maximum of 20) on the LSS bus before it gets interrupted. This allows it to insert arbitrary delays in a transfer at a byte boundary. For example the CPU may want to transmit 30 bytes to a QA chip but insert a delay between the 20^(th) and 21^(st) bytes sent. It does this by first writing 20 bytes to the data buffer. It then writes a command to generate a START condition, send the primary id byte and then transmit the 20 bytes from the data buffer. When interrupted by the LSS block to indicate successful completion of the command the CPU can then write the remaining 10 bytes to the data buffer. It can then wait for a defined period of time before writing a command to transmit the 10 bytes from the data buffer and generate a STOP condition to terminate the transaction over the LSS bus.

An interrupt to the CPU is generated for one cycle when any bit in LssNIntStatus is set. The CPU can read LssNIntStatus to discover the source of the interrupt. The LssNIntStatus registers are cleared when the CPU writes to the LssNCmd register. A null command write to the LssNCmd register will cause the LssNIntStatus registers to clear and no new command to start. A null command is defined as Start, IdbyteEnable, RdWrEnable and Stop all set to zero.

TABLE 104 LSS command register description bit(s) name Description 0 Start When 1, Issue a START condition on the LSS bus. 1 IdByteEnable ID byte transmit enable: 1 - transmit byte in IdByte field 0 - ignore byte in IdByte field 2 RdWrEnable Read/write transfer enable: 0 - ignore settings of RdWrSense, ReadNack and TxRxByteCount 1 - if RdWrSense is 0, then perform a write transfer of TxRxByteCount bytes from the data buffer. if RdWrSense is 1, then perform a read transfer of TxRxByteCount bytes into the data buffer. Each byte should be acknowl- edged and the last byte received is acknowledged/not-acknowledged according to the setting of ReadNack. 3 RdWrSense Read/write sense indicator: 0 - write 1 - read 4 ReadNack Indicates, for a read transfer, whether to issue an acknowledge or a not-acknowledge after the last byte received (indicated by TxRxByteCount). 0 - issue acknowledge after last byte received 1 - issue not-acknowledge after last byte received. 5 Stop When 1, issue a STOP condition on the LSS bus. 7:6 reserved Must be 0 15:8  IdByte Byte to be transmitted if IdByteEnable is 1. Bit 8 corresponds to the LSB. 20:16 TxRxByteCount Number of bytes to be transmitted from the data buffer or the number of bytes to be received into the data buffer. The maximum value that should be pro- grammed is 20, as the size of the data buffer is 20 bytes. Valid values are 1 to 20, 0 is valid when RdWrEnable = 0, other cases are invalid and un- defined.

The data buffer is implemented in the LSS master block. When the CPU writes to the LssNBuffer registers the data written is presented to the LSS master block via the lssN_buffer_wrdata bus and configuration registers block pulses the lssN_buffer_wen bit corresponding to the register written. For example if LssNBuffer[2] is written to lssN_buffer_wen[2] will be pulsed. When the CPU reads the LssNBuffer registers the configuration registers block reflect the lssN_buffer_rdata bus back to the CPU.

19.3.3 LSS Master Unit

The LSS master unit is instantiated for both LSS bus 0 and LSS bus 1. It controls transactions on the LSS bus by means of the state machine shown in FIG. 83, which interprets the commands that are written by the CPU. It also contains a single 20 byte data buffer used for transmitting and receiving data.

The CPU can write data to be transmitted on the LSS bus by writing to the LssNBuffer registers. It can also read data that the LSS master unit receives on the LSS bus by reading the same registers. The LSS master always transmits or receives bytes to or from the data buffer in the same order.

For a transmit command, LssNBuffer[0][7:0] gets transmitted first, then LssNBuffer[0][15:8], LssNBuffer[0][23:16], LssNBuffer[0][31:24], LssNBuffer[1][7:0] and so on until TxRxByteCount number of bytes are transmitted. A receive command fills data to the buffer in the same order. Each new command the buffer start point is reset.

All state machine outputs, flags and counters are cleared on reset. After a reset the state machine goes to the Reset state and initialises the LSS pins (lss_clk is set to 1, lss_data is tristated and allowed to be pulled up to 1). When the reset condition is removed the state machine transitions to the Wait state.

It remains in the Wait state until lss_new_cmd equals 1. If the Start bit of the command is 0 the state machine proceeds directly to the CheckIdByteEnable state. If the Start bit is 1 it proceeds to the GenerateStart state and issues a START condition on the LSS bus.

In the CheckIdByteEnable state, if the IdByteEnable bit of the command is 0 the state machine proceeds directly to the CheckRdWrEnable state. If the IdByteEnable bit is 1 the state machine enters the SendIdByte state and the byte in the IdByte field of the command is transmitted on the LSS. The WaitForIdAck state is then entered. If the byte is acknowledged, the state machine proceeds to the CheckRdWrEnable state. If the byte is not-acknowledged, the state machine proceeds to the GenerateInterrupt state and issues an interrupt to indicate a not-acknowledge was received after transmission of the primary id byte.

In the CheckRdWrEnable state, if the RdWrEnable bit of the command is 0 the state machine proceeds directly to the CheckStop state. If the RdWrEnable bit is 1, count is loaded with the value of the TxRxByteCount field of the command and the state machine enters either the ReceiveByte state if the RdWrSense bit of the command is 1 or the TransmitByte state if the RdWrSense bit is 0.

For a write transaction, the state machine keeps transmitting bytes from the data buffer, decrementing count after each byte transmitted, until count is 1. If all the bytes are successfully transmitted the state machine proceeds to the CheckStop state. If the slave QA chip not-acknowledges a transmitted byte, the state machine indicates this error by issuing an interrupt to the CPU and then entering the GenerateInterrupt state.

For a read transaction, the state machine keeps receiving bytes into the data buffer, decrementing count after each byte transmitted, until count is 1. After each byte received the LSS master must issue an acknowledge. After the last expected byte (i.e. when count is 1) the state machine checks the ReadNack bit of the command to see whether it must issue an acknowledge or not-acknowledge for that byte. The CheckStop state is then entered.

In the CheckStop state, if the Stop bit of the command is 0 the state machine proceeds directly to the GenerateInterrupt state. If the Stop bit is 1 it proceeds to the GenerateStop state and issues a STOP condition on the LSS bus before proceeding to the GenerateInterrupt state. In both cases an interrupt is issued to indicate successful completion of the command.

The state machine then enters the Wait state to await the next command. When the state machine reenters the Wait state the output pins (lss_and lss_clk) are not changed, they retain the state of the last command. This allows the possibility of multi-command transactions. The CPU may abort the current transfer at any time by performing a write to the Reset register of the LSS block.

19.3.3.1 START and STOP Generation

START and STOP conditions, which signal the beginning and end of data transmission, occur when the LSS master generates a falling and rising edge respectively on the data while the clock is high.

In the GenerateStart state, lss_gpio_clk is held high with lss_gpio_e remaining deasserted (so the data line is pulled high externally) for LssClockHighLowDuration pclk cycles. Then lss_gpio_e is asserted and lss_gpio_dout is pulled low (to drive a 0 on the data line, creating a falling edge) with lss_gpio_clk remaining high for another LssClockHighLowDuration pclk cycles. In the GenerateStop state, both lss_gpio_clk and lss_gpio_dout are pulled low followed by the assertion of lss_gpio_e to drive a 0 while the clock is low. After LssClockHighLowDuration pclk cycles, lss_gpio_clk is set high. After a further LssClockHighLowDuration pclk cycles, lss_gpio_e is deasserted to release the data bus and create a rising edge on the data bus during the high period of the clock.

If the bus is not in the required state for start and stop generation (lss_clk=1, lss_data=1 for start, and lss_clk=1, lss_data=0), the state machine moves the bus to the correct state and proceeds as described above. FIG. 82 shows the transition timing from any bus state to start and stop generation

19.3.3.2 Clock Pulse Generation

The LSS master holds lss_gpio_clk high while the LSS bus is inactive. A clock pulse is generated for each bit transmitted or received over the LSS bus. It is generated by first holding lss_gpio_clk low for LssClockHighLowDuration pclk cycles, and then high for LssClockHighLowDuration pclk cycles.

19.3.3.3 Data De-Glitching

When data is received in the LSS block it is passed to a de-glitching circuit. The de-glitch circuit samples the data 3 times on pclk and compares the samples. If all 3 samples are the same then the data is passed, otherwise the data is ignored.

Note that the LSS data input on SoPEC is double registered in the GPIO block before being passed to the LSS.

19.3.3.4 Data Reception

The input data, gpio_lss_di, is first synchronised to the pclk domain by means of two flip-flops clocked by pclk (the double register resides in the GPIO block). The LSS master generates a clock pulse for each bit received. The output lss_gpio_e is deasserted LssClockToDataHold pclk cycles after the falling edge of lss_gpio_clk to release the data bus. The value on the synchronised gpio_lss_di is sampled Tstrobe number of clock cycles after the rising edge of lss_gpio_clk (the data is de-glitched over a further 3 stage register to avoid possible glitch detection). See FIG. 84 for further timing information.

In the ReceiveByte state, the state machine generates 8 clock pulses. At each Tstrobe time after the rising edge of lss_gpio_clk the synchronised gpio_lss_di is sampled. The first bit sampled is LssNBuffer[0][7], the second LssNBuffer[0][6], etc to LssNBuffer[0][0]. For each byte received the state machine either sends an NAK or an ACK depending on the command configuration and the number of bytes received.

In the SendNack state the state machine generates a single clock pulse. lss_gpio_e is deasserted and the LSS data line is pulled high externally to issue a not-acknowledge.

In the SendAck state the state machine generates a single clock pulse. lss_gpio_e is asserted and a 0 driven on lss_gpio_dout after lss_gpio_clk falling edge to issue an acknowledge.

19.3.3.5 Data Transmission

The LSS master generates a clock pulse for each bit transmitted. Data is output on the LSS bus on the falling edge of lss_gpio_clk.

When the LSS master drives a logical zero on the bus it will assert lss_gpio_e and drive a 0 on lss_gpio_dout after lss_gpio_clk falling edge. lss_gpio_e will remain asserted and lss_gpio_dout will remain low until the next lss_clk falling edge.

When the LSS master drives a logical one lss_gpio_e should be deasserted at lss_gpio_clk falling edge and remain deasserted at least until the next lss_gpio_clk falling edge. This is because the LSS bus will be externally pulled up to logical one via a pull-up resistor.

In the SendId byte state, the state machine generates 8 clock pulses to transmit the byte in the IdByte field of the current valid command. On each falling edge of lss_gpio_clk a bit is driven on the data bus as outlined above. On the first falling edge IdByte[7] is driven on the data bus, on the second falling edge IdByte[6] is driven out, etc.

In the TransmitByte state, the state machine generates 8 clock pulses to transmit the byte at the output of the transmit FIFO. On each falling edge of lss_gpio_clk a bit is driven on the data bus as outlined above. On the first falling edge LssNBuffer[0][7] is driven on the data bus, on the second falling edge LssNBuffer[0][6] is driven out, etc on to LssNBuffer[0][7] bits.

In the WaitForAck state, the state machine generates a single clock pulse. At Tstrobe time after the rising edge of lss_gpio_clk the synchronized gpio_lss_di is sampled. A 0 indicates an acknowledge and ack_detect is pulsed, a 1 indicates a not-acknowledge and nack_detect is pulsed.

19.3.3.6 Data Rate Control

The CPU can control the data rate by setting the clock period of the LSS bus clock by programming appropriate value in LssClockHighLowDuration. The default setting for the register is 200 (pclk cycles) which corresponds to transmission rate of 400 kHz on the LSS bus (the lss_clk is high for LssClockHighLowDuration cycles then low for LssClockHighLowDuration cycles). The lss_clk will always have a 50:50 duty cycle. The LssClockHighLowDuration register should not be set to values less than 8.

The hold time of lss_data after the falling edge of lss_clk is programmable by the LssClocktoDataHold register. This register should not be programmed to less than 2 or greater than the LssClockHighLowDuration value.

19.3.3.7 LSS Master Timing Parameters

The LSS master timing parameters are shown in FIG. 84 and the associated values are shown in Table 105.

TABLE 105 LSS master timing parameters Parameter Description min nom max unit LSS Master Driving Tp LSS clock period divided by 2 8 200 FFFF pclk cycles Tstart_delay Time to start data edge from rising Tp + pclk cycles clock edge LssClocktoDataHold Tstop_delay Time to stop data edge from rising Tp + pclk cycles clock edge LssClocktoDataHold Tdata_setup Time from data setup to rising clock Tp − 2 − pclk cycles edge LssClocktoDataHold Tdata_hold Time from falling clock edge to data LssClocktoDataHold pclk cycles hold Tack_setup Time that outgoing (N)Ack is setup Tp − 2 − pclk cycles before lss_clk rising edge LssClocktoDataHold Tack_hold Time that outgoing (N)Ack is held LssClocktoDataHold pclk cycles after lss_clk falling edge LSS Master Sampling Tstrobe LSS master strobe point for Tp − 2 Tp − 2 pclk cycles incoming data and (N)Ack values

DRAM Subsystem

20 DRAM Interface Unit (DIU)

20.1 Overview

FIG. 85 shows how the DIU provides the interface between the on-chip 20 Mbit embedded DRAM and the rest of SoPEC. In addition to outlining the functionality of the DIU, this chapter provides a top-level overview of the memory storage and access patterns of SoPEC and the buffering required in the various SoPEC blocks to support those access requirements.

The main functionality of the DIU is to arbitrate between requests for access to the embedded DRAM and provide read or write accesses to the requesters. The DIU must also implement the initialisation sequence and refresh logic for the embedded DRAM.

The arbitration scheme uses a fully programmable timeslot mechanism for non-CPU requesters to meet the bandwidth and latency requirements for each unit, with unused slots re-allocated to provide best effort accesses. The CPU is allowed high priority access, giving it minimum latency, but allowing bounds to be placed on its bandwidth consumption.

The interface between the DIU and the SoPEC requesters is similar to the interface on PEC1 i.e. separate control, read data and write data busses.

The embedded DRAM is used principally to store:

-   -   CPU program code and data.     -   PEP (re)programming commands.     -   Compressed pages containing contone, bi-level and raw tag data         and header information.     -   Decompressed contone and bi-level data.     -   Dotline store during a print.     -   Print setup information such as tag format structures, dither         matrices and dead nozzle information.         20.2 IBM Cu-11 Embedded DRAM         20.2.1 Single Bank

SoPEC will use the 1.5 V core voltage option in IBM's 0.13 μm class Cu-11 process.

The random read/write cycle time and the refresh cycle time is 3 cycles at 160 MHz [16]. An open page access will complete in 1 cycle if the page mode select signal is clocked at 320 MHz or 2 cycles if the page mode select signal is clocked every 160 MHz cycle. The page mode select signal will be clocked at 160 MHz in SoPEC in order to simplify timing closure. The DRAM word size is 256 bits.

Most SoPEC requesters will make single 256 bit DRAM accesses (see Section 20.4). These accesses will take 3 cycles as they are random accesses i.e. they will most likely be to a different memory row than the previous access.

The entire 20 Mbit DRAM will be implemented as a single memory bank. In Cu-11, the maximum single instance size is 16 Mbit. The first 1 Mbit tile of each instance contains an area overhead so the cheapest solution in terms of area is to have only 2 instances. 16 Mbit and 4 Mbit instances would together consume an area of 14.63 mm² as would 2 times 10 Mbit instances. 4 times 5 Mbit instances would require 17.2 mm².

The instance size will determine the frequency of refresh. Each refresh requires 3 clock cycles. In Cu-11 each row consists of 8 columns of 256-bit words. This means that 10 Mbit requires 5120 rows. A complete DRAM refresh is required every 3.2 ms. Two times 10 Mbit instances would require a refresh every 100 clock cycles, if the instances are refreshed in parallel.

The SoPEC DRAM will be constructed as two 10 Mbit instances implemented as a single memory bank.

20.3 SoPEC Memory Usage Requirements

The memory usage requirements for the embedded DRAM are shown in Table 106.

TABLE 106 Memory Usage Requirements Block Size Description Compressed 2048 Kbytes Compressed data page store for page Bi-level and contone data store Decompressed 108 Kbyte 13824 lines with scale factor Contone 6 = 2304 pixels, store 12 Store lines, 4 colors = 108 kB 13824 lines with scale factor 5 = 2765 pixels, store 12 lines, 4 colors = 130 kB Spot line 5.1 Kbyte 13824 dots/line so 3 lines store is 5.1 kB Tag Format Typically 55 kB in for 384 dot line tags Structure 12 Kbyte 2.5 mm tags (1/10th inch) @ (2.5 mm tags 1600 dpi require 160 dot lines = @ 800 dpi) 160/384 × 55 or 23 kB 2.5 mm tags (1/10th inch) @ 800 dpi require 80/384 × 55 = 12 kB Dither Matrix 4 Kbytes 64 × 64 dither matrix is 4 kB store 128 × 128 dither matrix is 16 kB 256 × 256 dither matrix is 64 kB DNC Dead 1.4 Kbytes Delta encoded, (10 bit delta Nozzle Table position + 6 dead nozzle mask) × % Dnozzle 5% dead nozzles requires (10 + 6) × 692 Dnozzles = 1.4 Kbytes Dot-line 369.6 Kbytes Assume each color row is sepa- store rated by 5 dot lines on the print head The dot line store will be 0 + 5 + 10 . . . 50 + 55 = 330 half dot lines + 48 extra half dot lines (4 per dot row) + 60 extra half dot lines estimated to account for printhead mis- alignment = 438 half dot lines. 438 half dot lines of 6912 dots = 369.6 Kbytes PCU Program 8 Kbytes 1024 commands of 64 bits = code 8 KB CPU 64 Kbytes Program code and data TOTAL 2620 Kbytes (12 Kbyte TFS storage) Note: Total storage is fixed to 2560 Kbytes to align to 20 Mbit DRAM. This will mean that less space than noted in Table may be available for the compressed band store. 20.4 SoPEC Memory Access Patterns

Table 107 shows a summary of the blocks on SoPEC requiring access to the embedded DRAM and their individual memory access patterns. Most blocks will access the DRAM in single 256-bit accesses. All accesses must be padded to 256-bits except for 64-bit CDU write accesses and CPU write accesses. Bits which should not be written are masked using the individual DRAM bit write inputs or byte write inputs, depending on the foundry. Using single 256-bit accesses means that the buffering required in the SoPEC DRAM requesters will be minimized.

TABLE 107 Memory access patterns of SoPEC DRAM Requesters DRAM requester Direction Memory access pattern CPU R Single 256-bit reads. W Single 32-bit, 16-bit or 8-bit writes. SCB R Single 256-bit reads. W Single 256-bit writes, with byte enables. CDU R Single 256-bit reads of the compressed contone data. W Each CDU access is a write to 4 consecutive DRAM words in the same row but only 64 bits of each word are written with the remaining bits write masked. The access time for this 4 word page mode burst is 3 + 2 + 2 + 2 = 9 cycles if the page mode select signal is clocked at 160 MHz. CFU R Single 256 bit reads. LBD R Single 256 bit reads. SFU R Separate single 256 bit reads for previous and current line but sharing the same DIU interface W Single 256 bit writes. TE(TD) R Single 256 bit reads. Each read returns 2 times 128 bit tags. TE(TFS) R Single 256 bit reads. TFS is 136 bytes. This means there is unused data in the fifth 256 bit read. A total of 5 reads is required. HCU R Single 256 bit reads. 128 × 128 dither matrix requires 4 reads per line with double buffering. 256 × 256 dither matrix requires 8 reads at the end of the line with single buffering. DNC R Single 256 bit dead nozzle table reads. Each dead nozzle table read contains 16 dead-nozzle tables entries each of 10 delta bits plus 6 dead nozzle mask bits. DWU W Single 256 bit writes since enable/disable DRAM access per color plane. LLU R Single 256 bit reads since enable/disable DRAM access per color plane. PCU R Single 256 bit reads. Each PCU command is 64 bits so each 256 bit word can contain 4 PCU commands. PCU reads from DRAM used for repro- gramming PEP should be executed with minimum latency. If this occurs between pages then there will be free bandwidth as most of the other SoPEC Units will not be requesting from DRAM. If this occurs between bands then the LDB, CDU and TE bandwidth will be free. So the PCU should have a high priority to access to any spare bandwidth. Refresh Single refresh. 20.5 Buffering Required in SoPEC DRAM Requesters

If each DIU access is a single 256-bit access then we need to provide a 256-bit double buffer in the DRAM requester. If the DRAM requester has a 64-bit interface then this can be implemented as an 8×64-bit FIFO.

TABLE 108 Buffer sizes in SoPEC DRAM requesters DRAM Direc- Buffering required Requester tion Access patterns in block CPU R Single 256-bit reads. Cache. W Single 32-bit writes None. but allowing 16-bit or byte addressable writes. SCB R Single 256-bit reads. Double 256-bit buffer. W Single 256-bit writes, Double 256-bit with byte enables. buffer. CDU R Single 256-bit reads Double 256-bit of the compressed buffer. contone data. W Each CDU access is a Double half JPEG write to 4 consecutive block buffer. DRAM words in the same row but only 64 bits of each word are written with the remaining bits write masked. CFU R Single 256 bit reads. Triple 256-bit buffer. LBD R Single 256 bit reads. Double 256-bit buffer. SFU R Separate single 256 bit Double 256-bit reads for previous and buffer for each current line but sharing read channel. the same DIU interface W Single 256 bit writes. Double 256-bit buffer. TE(TD) R Single 256 bit reads. Double 256-bit buffer. TE(TFS) R Single 256 bit reads. Double line-buffer TFS is 136 bytes. for 136 bytes This means there is implemented in TE. unused data in the fifth 256 bit read. A total of 5 reads is required. HCU R Single 256 bit reads. Configurable 128 × 128 dither between double matrix requires 4 reads 128 byte buffer per line with double and single 256 buffering. 256 × 256 byte buffer. dither matrix requires 8 reads at the end of the line with single buffering. DNC R Single 256 bit reads Double 256-bit buffer. Deeper buffering could be specified to cope with local clusters of dead nozzles. DWU W Single 256 bit writes Double 256-bit per enabled odd/even buffer per color color plane. plane. LLU R Single 256 bit reads Double 256-bit per enabled odd/even buffer per color color plane. plane. PCU R Single 256 bit reads. Single 256-bit Each PCU command is buffer. 64 bits so each 256 bit DRAM read can contain 4 PCU commands. Requested command is read from DRAM together with the next 3 contiguous 64-bits which are cached to avoid unnecessary DRAM reads. Refresh Single refresh. None. 20.6 SoPEC DIU Bandwidth Requirements

TABLE 109 SoPEC DIU Bandwidth Requirements Number of cycles between Peak each 256-bit Bandwidth Example DRAM access which must Average number of Block to meet peak be supplied Bandwidth allocated Name Direction bandwidth bits/cycle) (bits/cycle) timeslots¹ CPU R W SCB R W 3482 0.734 0.3933 1 CDU R 128 (SF = 4), 64/n2 (SF = n), 32/10*n2 1 (SF = 6) 288 (SF = 6), 1.8 (SF = 6), (SF = n), 2 (SF = 4) 1:1 compression4 4 (SF = 4) 0.09 (SF = 6), (1:1 compression) 0.2 (SF = 4) (10:1 com- pression)5 W For individual 64/n2 (SF = n), 32/n2 (SF = n)7, 2 (SF = 6)8 accesses: 16 1.8 (SF = 6), 0.9 (SF = 6), 4 (SF = 4) cycles (SF = 4), 4 (SF = 4) 2 (SF = 4) 36 cycles (SF = 6), n2 cycles (SF = n). Will be implemented as a page mode burst of 4 accesses every 64 cycles (SF = 4), 144 (SF = 6), 4*n2 (SF = n) cycles6 CFU R 32 (SF = 4), 32/n (SF = n), 32/n (SF = n), 6 (SF = 6) 48 (SF = 6)9 5.4 (SF = 6), 5.4 (SF = 6), 8 (SF = 4) 8 (SF = 4) 8 (SF = 4) LBD R 256 (1:1 com- 1 (1:1 com- 0.1 (10:1 com- 1 pression)10 pression) pression)11 SFU R 12812 2 2 2 W 25613 1 1 1 TE(TD) R 25214 1.02 1.02 1 TE(TFS) R 5 reads per line 15 0.093 0.093 0 HCU R 4 reads per line 0.074 0.074 0 for 128 × 128 dither matrix16 DNC R 106 (5% dead- 2.4 (clump of 0.8 (equally 3 nozzles 10-bit dead nozzles) spaced dead delta encoded)17 nozzles) DWU W 6 writes every 6 6 6 25618 LLU R 8 reads every 8 6 8 25619 PCU R 25620 1 1 1 Refresh 10021 2.56 2.56 3 (effective) TOTAL SF = 6:34.9 SF = 6:27.5 SF = 6:36 SF = 4:41.9 SF = 4:31.2 excluding CPU. excluding CPU excluding CPU SF = 4:41 excluding CPU Notes: 1: The number of allocates timeslots is based on 64 timeslots each of 1 bit/cycle but broken down to a granularity of 0.25 bit/cycle. Bandwidth is allocated based on peak bandwidth. 2: Wire-speed bandwidth for a 4 wire SCB configuration is 32 Mbits/s for each wire plus 12 Mbits/s for USB. This is a maximum of 138 Mbits/s. The maximum effective data rate is 26 Mbits/s for each wire plus 8 Mbits/s for USB. This is 112 Mbits/s. 112 Mbits/s is 0.734 bits/cycle or 256 bits every 348 cycles. 3: Wire-speed bandwidth for a 2 wire SCB configuration is 32 Mbits/s for each wire plus 12 Mbits/s for USB. This is a maximum of 74 Mbits/s. The maximum effective data rate is 26 Mbits/s for each wire plus 8 Mbits/s for USB. This is 60 Mbits/s. 60 Mbits/s is 0.393 bits/cycle or 256 bits every 650 cycles. 4: At 1:1 compression CDU must read 4 color pixel (32 bits) every SF² cycles. 5: At 10:1 average compression CDU must read 4 color pixel (32 bits) every 10*SF² cycles. 6: 4 color pixel (32 bits) is required, on average, by the CFU every SF² (scale factor) cycles. The time available to write the data is a function of the size of the buffer in DRAM. 1.5 buffering means 4 color pixel (32 bits) must be written every SF²/2 (scale factor) cycles. Therefore, at a scale factor of SF, 64 bits are required every SF² cycles. Since 64 valid bits are written per 256-bit write (Figure on page 379 on page Error! Bookmark not defined.) then the DRAM is accessed every SF² cycles i.e. at SF4 an access every 16 cycles, at SF6 an access every 36 cycles. If a page mode burst of 4 accesses is used then each access takes (3 + 2 + 2 + 2) equals 9 cycles. This means at SF, a set of 4 back-to-back accesses must occur every 4*SF² cycles. This assumes the page mode select signal is clocked at 160 MHz. CDU timeslots therefore take 9 cycles. For scale factors lower than 4 double buffering will be used. 7: The peak bandwidth is twice the average bandwidth in the case of 1.5 buffering. 8: Each CDU(W) burst takes 9 cycles instead of 4 cycles for other accesses so CDU timeslots are longer. 9: 4 color pixel (32 bits) read by CFU every SF cycles. At SF4, 32 bits is required every 4 cycles or 256 bits every 32 cycles. At SF6, 32bits every 6 cycles or 256 bits every 48 cycles. 10: At 1:1 compression require 1 bit/cycle or 256 bits every 256 cycles. 11: The average bandwidth required at 10:1 compression is 0.1 bits/cycle. 12: Two separate reads of 1 bit/cycle. 13: Write at 1 bit/cycle. 14: Each tag can be consumed in at most 126 dot cycles and requires 128 bits. This is a maximum rate of 256 bits every 252 cycles. 15: 17 × 64 bit reads per line in PEC1 is 5 × 256 bit reads per line in SoPEC. Double-line buffered storage. 16: 128 bytes read per line is 4 × 256 bit reads per line. Double-line buffered storage. 17: 5% dead nozzles 10-bit delta encoded stored with 6-bit dead nozzle mask requires 0.8 bits/cycle read access or a 256-bit access every 320 cycles. This assumes the dead nozzles are evenly spaced out. In practice dead nozzles are likely to be clumped. Peak bandwidth is estimated as 3 times average bandwidth. 18: 6 bits/cycle requires 6 × 256 bit writes every 256 cycles. 19: 6 bits/160 MHz SoPEC cycle average but will peak at 2 × 6 bits per 106 MHz print head cycle or 8 bits/SoPEC cycle. The PHI can equalise the DRAM access rate over the line so that the peak rate equals the average rate of 6 bits/cycle. The print head is clocked at an effective speed of 106 MHz. 20: Assume one 256 read per 256 cycles is sufficient i.e. maximum latency of 256 cycles per access is allowable. 21: Refresh must occur every 3.2 ms. Refresh occurs row at a time over 5120 rows of 2 parallel 10 Mbit instances. Refresh must occur every 100 cycles. Each refresh takes 3 cycles. 20.7 DIU BUS Topology 20.7.1 Basic Topology

TABLE 110 SoPEC DIU Requesters Read Write Other CPU CPU Refresh SCB SCB CDU CDU CFU SFU LBD DWU SFU TE(TD) TE(TFS) HCU DNC LLU PCU

Table 110 shows the DIU requesters in SoPEC. There are 12 read requesters and 5 write requesters in SoPEC as compared with 8 read requesters and 4 write requesters in PEC1. Refresh is an additional requester.

In PEC1, the interface between the DIU and the DIU requesters had the following main features:

-   -   separate control and address signals per DIU requester         multiplexed in the DIU according to the arbitration scheme,     -   separate 64-bit write data bus for each DRAM write requester         multiplexed in the DIU,     -   common 64-bit read bus from the DIU with separate enables to         each DIU read requester.

Timing closure for this bussing scheme was straight-forward in PEC1. This suggests that a similar scheme will also achieve timing closure in SoPEC. SoPEC has 5 more DRAM requesters but it will be in a 0.13 um process with more metal layers and SoPEC will run at approximately the same speed as PEC1.

Using 256-bit busses would match the data width of the embedded DRAM but such large busses may result in an increase in size of the DIU and the entire SoPEC chip. The SoPEC requestors would require double 256-bit wide buffers to match the 256-bit busses. These buffers, which must be implemented in flip-flops, are less area efficient than 8-deep 64-bit wide register arrays which can be used with 64-bit busses. SoPEC will therefore use 64-bit data busses. Use of 256-bit busses would however simplify the DIU implementation as local buffering of 256-bit DRAM data would not be required within the DIU.

20.7.1.1 CPU DRAM Access

The CPU is the only DIU requestor for which access latency is critical. All DIU write requesters transfer write data to the DIU using separate point-to-point busses. The CPU will use the cpu_dataout[31:0] bus. CPU reads will not be over the shared 64-bit read bus. Instead, CPU reads will use a separate 256-bit read bus.

20.7.2 Making More Efficient Use of DRAM Bandwidth

The embedded DRAM is 256-bits wide. The 4 cycles it takes to transfer the 256-bits over the 64-bit data busses of SoPEC means that effectively each access will be at least 4 cycles long. It takes only 3 cycles to actually do a 256-bit random DRAM access in the case of IBM DRAM.

20.7.2.1 Common Read Bus

If we have a common read data bus, as in PEC1, then if we are doing back to back read accesses the next DRAM read cannot start until the read data bus is free. So each DRAM read access can occur only every 4 cycles. This is shown in FIG. 86 with the actual DRAM access taking 3 cycles leaving 1 unused cycle per access.

20.7.2.2 Interleaving CPU and Non-CPU Read Accesses

The CPU has a separate 256-bit read bus. All other read accesses are 256-bit accesses are over a shared 64-bit read bus. Interleaving CPU and non-CPU read accesses means the effective duration of an interleaved access timeslot is the DRAM access time (3 cycles) rather than 4 cycles.

FIG. 87 shows interleaved CPU and non-CPU read accesses.

20.7.2.3 Interleaving Read and Write Accesses

Having separate write data busses means write accesses can be interleaved with each other and with read accesses. So now the effective duration of an interleaved access timeslot is the DRAM access time (3 cycles) rather than 4 cycles. Interleaving is achieved by ordering the DIU arbitration slot allocation appropriately.

FIG. 88 shows interleaved read and write accesses. FIG. 89 shows interleaved write accesses.

256-bit write data takes 4 cycles to transmit over 64-bit busses so a 256-bit buffer is required in the DIU to gather the write data from the write requester. The exception is CPU write data which is transferred in a single cycle.

FIG. 89 shows multiple write accesses being interleaved to obtain 3 cycle DRAM access. Since two write accesses can overlap two sets of 256-bit write buffers and multiplexors to connect two write requestors simultaneously to the DIU are required.

Write requestors only require approximately one third of the total non-CPU bandwidth. This means that a rule can be introduced such that non-CPU write requestors are not allocated adjacent timeslots. This means that a single 256-bit write buffer and multiplexor to connect the one write requestor at a time to the DIU is all that is required.

Note that if the rule prohibiting back-to-back non-CPU writes is not adhered to, then the second write slot of any attempted such pair will be disregarded and re-allocated under the unused read round-robin scheme.

20.7.3 Bus Widths Summary

TABLE 111 SoPEC DIU Requesters Data Bus Width Bus access Read Bus access width Write width CPU 256 (separate) CPU 32 SCB 64 (shared) SCB 64 CDU 64 (shared) CDU 64 CFU 64 (shared) SFU 64 LBD 64 (shared) DWU 64 SFU 64 (shared) TE(TD) 64 (shared) TE(TFS) 64 (shared) HCU 64 (shared) DNC 64 (shared) LLU 64 (shared) PCU 64 (shared) 20.7.4 Conclusions

Timeslots should be programmed to maximise interleaving of shared read bus accesses with other accesses for 3 cycle DRAM access. The interleaving is achieved by ordering the DIU arbitration slot allocation appropriately. CPU arbitration has been designed to maximise interleaving with non-CPU requesters

20.8 SoPEC DRAM Addressing Scheme

The embedded DRAM is composed of 256-bit words. However the CPU-subsystem may need to write individual bytes of DRAM. Therefore it was decided to make the DIU byte addressable. 22 bits are required to byte address 20 Mbit of DRAM.

Most blocks read or write 256 bit words of DRAM. Therefore only the top 17 bits i.e. bits 21 to 5 are required to address 256-bit word aligned locations.

The exceptions are

-   -   CDU which can write 64-bits so only the top 19 address bits i.e.         bits 21-3 are required.     -   CPU writes can be 8, 16 or 32-bits. The cpu_diu_wmask[1:0] pins         indicate whether to write 8, 16 or 32 bits.

All DIU accesses must be within the same 256-bit aligned DRAM word. The exception is the CDU write access which is a write of 64-bits to each of 4 contiguous 256-bit DRAM words.

20.8.1 Write Address Constaints Specific to the CDU

Note the following conditions which apply to the CDU write address, due to the four masked page-mode writes which occur whenever a CDU write slot is arbitrated.

-   -   The CDU address presented to the DIU is cdu_diu_wadr[21:3].     -   Bits [4:3] indicate which 64-bit segment out of 256 bits should         be written in 4 successive masked page-mode writes.     -   Each 10-Mbit DRAM macro has an input address port of width         [15:0]. Of these bits, [2:0] are the “page address”. Page-mode         writes, where you just vary these LSBs (i.e. the “page” or         column address), but keep the rest of the address constant, are         faster than random writes. This is taken advantage of for CDU         writes.     -   To guarantee against trying to span a page boundary, the DIU         treats “cdu_diu_wadr[6:5]” as being fixed at “00”.     -   From cdu_diu_wadr[21:3], a initial address of         cdu_diu_wadr[21:7], concatenated with “00”, is used as the         starting location for the first CDU write. This address is then         auto-incremented a further three times.         20.9 DIU Protocols

The DIU protocols are

-   -   Pipelined i.e. the following transaction is initiated while the         previous transfer is in progress.     -   Split transaction i.e. the transaction is split into independent         address and data transfers.         20.9.1 Read Protocol Except CPU

The SoPEC read requestors, except for the CPU, perform single 256-bit read accesses with the read data being transferred from the DIU in 4 consecutive cycles over a shared 64-bit read bus, diu_data[63:0]. The read address <unit>_diu_radr[21:5] is 256-bit aligned.

The read protocol is:

-   -   <unit>_diu_rreq is asserted along with a valid         <unit>_diu_radr[21:5].     -   The DIU acknowledges the request with diu_<unit>_rack. The         request should be deasserted. The minimum number of cycles         between <unit>_diu_rreq being asserted and the DIU generating an         diu_<unit>_rack strobe is 2 cycles (1 cycle to register the         request, 1 cycle to perform the arbitration—see Section         20.14.10).     -   The read data is returned on diu_data[63:0] and its validity is         indicated by diu_<unit>_rvalid. The overall 256 bits of data are         transferred over four cycles in the order:         [63:0]->[127:64]->[191:128]->[255:192].     -   When four diu_<unit>_rvalid pulses have been received then if         there is a further request <unit>_diu_rreq should be asserted         again. diu_<unit>_rvalid will be always be asserted by the DIU         for four consecutive cycles. There is a fixed gap of 2 cycles         between diu_<unit>_rack and the first diu_<unit>_rvalid pulse.         For more detail on the timing of such reads and the implications         for back-to-back sequences, see Section 20.14.10.         20.9.2 Read Protocol for CPU

The CPU performs single 256-bit read accesses with the read data being transferred from the DIU over a dedicated 256-bit read bus for DRAM data, dram_cpu_data[255.0]. The read address cpu_adr[21:5] is 256-bit aligned.

The CPU DIU read protocol is:

-   -   cpu_diu_rreq is asserted along with a valid cpu_adr[21:5].     -   The DIU acknowledges the request with diu_cpu_rack. The request         should be deasserted. The minimum number of cycles between         cpu_diu_rreq being asserted and the DIU generating a         cpu_diu_rack strobe is 1 cycle (1 cycle to perform the         arbitration—see Section 20.14.10).     -   The read data is returned on dram_cpu_data[255:0] and its         validity is indicated by diu_cpu_rvalid.     -   When the diu_cpu_rvalid pulse has been received then if there is         a further request cpu_diu_rreq should be asserted again. The         diu_cpu_rvalid pulse with a gap of 1 cycle after rack (1 cycle         for the read data to be returned from the DRAM—see Section         20.14.10).         20.9.3 Write Protocol Except CPU and CDU

The SoPEC write requestors, except for the CPU and CDU, perform single 256-bit write accesses with the write data being transferred to the DIU in 4 consecutive cycles over dedicated point-to-point 64-bit write data busses. The write address <unit>_diu_wadr[21:5] is 256-bit aligned.

The write protocol is:

-   -   <unit>_diu_wreq is asserted along with a valid         <unit>_diu_wadr[21:5].     -   The DIU acknowledges the request with diu_<unit>_wack. The         request should be deasserted. The minimum number of cycles         between <unit>_diu_wreq being asserted and the DIU generating an         diu_<unit>_wack strobe is 2 cycles (1 cycle to register the         request, 1 cycle to perform the arbitration—see Section         20.14.10).     -   In the clock cycles following diu_<unit>_wack the SoPEC Unit         outputs the 15<unit>_diu_data[63:0], asserting         <unit>_diu_wvalid. The first <unit>_diu_wvalid pulse can occur         the clock cycle after diu_<unit>_wack. <unit>_diu_wvalid remains         asserted for the following 3 clock cycles. This allows for         reading from an SRAM where new data is available in the clock         cycle after the address has changed e.g. the address for the         second 64-bits of write data is available the cycle after         diu_<unit>_wack meaning the second 64-bits of write data is a         further cycle later. The overall 256 bits of data is transferred         over four cycles in the order:         [63:0]->[127:64]->[191:128]->[255:192].     -   Note that for SCB writes, each 64-bit quarter-word has an 8-bit         byte enable mask associated with it. A different mask is used         with each quarter-word. The 4 mask values are transferred along         with their associated data, as shown in FIG. 92.     -   If four consecutive <unit>_diu_wvalid pulses are not provided by         the requester, then the arbitration logic will disregard the         write and re-allocate the slot under the unused read round-robin         scheme.

Once all the write data has been output then if there is a further request <unit>_diu_wreq should be asserted again.

20.9.4 CPU Write Protocol

The CPU performs single 128-bit writes to the DIU on a dedicated write bus, cpu_diu_wdata[127:0]. There is an accompanying write mask, cpu_diu_wmask[15:0], consisting of 16 byte enables and the CPU also supplies a 128-bit aligned write address on cpu_diu_wadr[21:4]. Note that writes are posted by the CPU to the DIU and stored in a 1-deep 35 buffer. When the DAU subsequently arbitrates in favour of the CPU, the contents of the buffer are written to DRAM.

The CPU write protocol, illustrated in FIG. 93, is as follows:—

-   -   The DIU signals to the CPU via diu_cpu_write_rdy that its write         buffer is empty and that the CPU may post a write whenever it         wishes.     -   The CPU asserts cpu_diu_wdatavalid to enable a write into the         buffer and to confirm the validity of the write address, data         and mask.     -   The DIU de-asserts diu_cpu_write_rdy in the following cycle to         indicate that its buffer is full and that the posted write is         pending execution.     -   When the CPU is next awarded a DRAM access by the DAU, the         buffer's contents are written to memory. The DIU re-asserts         diu_cpu_write_rdy once the write data has been captured by DRAM,         namely in the “MSN1” DCU state.     -   The CPU can then, if it wishes, asynchronously use the new value         of .diu_cpu_write_rdy to enable a new posted write in the same         “MSN1” cycle.         20.9.5 CDU Write Protocol

The CDU performs four 64-bit word writes to 4 contiguous 256-bit DRAM addresses with the first address specified by cdu_diu_wadr[21:3]. The write address cdu_diu_wadr[21:5] is 256-bit aligned with bits cdu_diu_wadr[4:3] allowing the 64-bit word to be selected.

The write protocol is:

-   -   cdu_diu_wdata is asserted along with a valid cdu_diu_wadr[21:3].     -   The DIU acknowledges the request with diu_cdu_wack. The request         should be deasserted. The minimum number of cycles between         cdu_diu_wreq being asserted and the DIU generating an         diu_cdu_wack strobe is 2 cycles (1 cycle to register the         request, 1 cycle to perform the arbitration—see Section         20.14.10).     -   In the clock cycles following diu_cdu_wack the CDU outputs the         cdu_diu_data[63:0], together with asserted cdu_diu_wvalid. The         first cdu_diu_wvalid pulse can occur the clock cycle after         diu_cdu_wack. cdu_diu_wvalid remains asserted for the following         3 clock cycles. This allows for reading from an SRAM where new         data is available in the clock cycle after the address has         changed e.g. the address for the second 64-bits of write data is         available the cycle after diu_cdu_wack meaning the second         64-bits of write data is a further cycle later. Data is         transferred over the 4-cycle window in an order, such that each         successive 64 bits will be written to a monotonically increasing         (by 1 location) 256-bit DRAM word.     -   If four consecutive cdu_diu_wvalid pulses are not provided with         the data, then the arbitration logic will disregard the write         and re-allocate the slot under the unused read round-robin         scheme.     -   Once all the write data has been output then if there is a         further request cdu_diu_wreq should be asserted again.         20.10 DIU Arbitration Mechanism

The DIU will arbitrate access to the embedded DRAM. The arbitration scheme is outlined in the next sections.

20.10.1 Timeslot Based Arbitration Scheme

Table summarised the bandwidth requirements of the SoPEC requestors to DRAM. If we allocate the DIU requestors in terms of peak bandwidth then we require 35.25 bits/cycle (at SF=6) and 40.75 bits/cycle (at SF=4) for all the requestors except the CPU.

A timeslot scheme is defined with 64 main timeslots. The number of used main timeslots is programmable between 1 and 64.

Since DRAM read requesters, except for the CPU, are connected to the DIU via a 64-bit data bus each 256-bit DRAM access requires 4 pclk cycles to transfer the read data over the shared read bus. The timeslot rotation period for 64 timeslots each of 4 pclk cycles is 256 pclk cycles or 1.6 μs, assuming pclk is 160 MHz. Each timeslot represents a 256-bit access every 256 pclk cycles or 1 bit/cycle. This is the granularity of the majority of DIU requesters bandwidth requirements in Table.

The SoPEC DIU requesters can be represented using 4 bits (Table n page 288 on page 268). Using 64 timeslots means that to allocate each timeslot to a requester, a total of 64×5-bit configuration registers are required for the 64 main timeslots.

Timeslot based arbitration works by having a pointer point to the current timeslot. When re-arbitration is signaled the arbitration winner is the current timeslot and the pointer advances to the next timeslot. Each timeslot denotes a single access. The duration of the timeslot depends on the access.

Note that advancement through the timeslot rotation is dependent on an enable bit, RotationSync, being set. The consequences of clearing and setting this bit are described in section 20.14.12.2.1 on page 295.

If the SoPEC Unit assigned to the current timeslot is not requesting then the unused timeslot arbitration mechanism outlined in Section 20.10.6 is used to select the arbitration winner.

Note that there is always an arbitration winner for every slot. This is because the unused read re-allocation scheme includes refresh in its round-robin protocol. If all other blocks are not requesting, an early refresh will act as fall-back for the slot.

20.10.2 Separate Read and Write Arbitration Windows

For write accesses, except the CPU, 256-bits of write data are transferred from the SoPEC DIU write requestors over 64-bit write busses in 4 clock cycles. This write data transfer latency means that writes accesses, except for CPU writes and also the CDU, must be arbitrated 4 cycles in advance. (The CDU is an exception because CDU writes can start once the first 64-bits of write data have been transferred since each 64-bits is associated with a write to a different 256-bit word).

Since write arbitration must occur 4 cycles in advance, and the minimum duration of a timeslot duration is 3 cycles, the arbitration rules must be modified to initiate write accesses in advance. Accordingly, there is a write timeslot lookahead pointer shown in FIG. 96 two timeslots in advance of the current timeslot pointer.

The following examples illustrate separate read and write timeslot arbitration with no adjacent write timeslots. (Recall rule on adjacent write timeslots introduced in Section 20.7.2.3 on page 238.)

In FIG. 97 writes are arbitrated two timeslots in advance. Reads are arbitrated in the same timeslot as they are issued. Writes can be arbitrated in the same timeslot as a read. During arbitration the command address of the arbitrated SoPEC Unit is captured.

Other examples are shown in FIG. 98 and FIG. 99. The actual timeslot order is always the same as the programmed timeslot order i.e. out of order accesses do not occur and data coherency is never an issue.

Each write must always incur a latency of two timeslots.

Startup latency may vary depending on the position of the first write timeslot. This startup latency is not important.

Table 112 shows the 4 scenarIOs depending on whether the current timeslot and write timeslot lookahead pointers point to read or write accesses.

TABLE 112 Arbitration with separate windows for read and write accesses current write timeslot timeslot lookahead pointer pointer actions Read write Initiate DRAM read, Initiate write arbitration Read1 read2 Initiate DRAM read1. Write1 write2 Initiate write2 arbitration. Execute DRAM write1. Write read Execute DRAM write.

If the current timeslot pointer points to a read access then this will be initiated immediately. If the write timeslot lookahead pointer points to a write access then this access is arbitrated immediately, or immediately after the read access associated with the current timeslot pointer is initiated.

When a write access is arbitrated the DIU will capture the write address. When the current timeslot pointer advances to the write timeslot then the actual DRAM access will be initiated. Writes will therefore be arbitrated 2 timeslots in advance of the DRAM write occurring.

At initialisation, the write lookahead pointer points to the first timeslot. The current timeslot pointer is invalid until the write lookahead pointer advances to the third timeslot when the current timeslot pointer will point to the first timeslot. Then both pointers advance in tandem.

CPU write accesses are excepted from the lookahead mechanism.

If the selected SoPEC Unit is not requesting then there will be separate read and write selection for unused timeslots. This is described in Section 20.10.6.

20.10.3 Arbitration of CPU Accesses

What distinguishes the CPU from other SoPEC requesters, is that the CPU requires minimum latency DRAM access i.e. preferably the CPU should get the next available timeslot whenever it requests.

The minimum CPU read access latency is estimated in Table 113. This is the time between the CPU making a request to the DIU and receiving the read data back from the DIU.

TABLE 113 Estimated CPU read access latency ignoring caching CPU read access latency Duration CPU cache miss 1 cycle CPU MMU logic issues request and 1 cycle DIU arbitration completes Transfer the read address to the DRAM 1 cycle DRAM read latency 1 cycle Register the read data in CPU bridge 1 cycle Register the read data in CPU 1 cycle CPU cache miss 1 cycle CPU MMU logic issues request and 1 cycle DIU arbitration completes TOTAL gap between requests 6 cycles

If the CPU, as is likely, requests DRAM access again immediately after receiving data from the DIU then the CPU could access every second timeslot if the access latency is 6 cycles. This assumes that interleaving is employed so that timeslots last 3 cycles. If the CPU access latency were 7 cycles, then the CPU would only be able to access every third timeslot.

If a cache hit occurs the CPU does not require DRAM access. For its next DIU access it will have to wait for its next assigned DIU slot. Cache hits therefore will reduce the number of DRAM accesses but not speed up any of those accesses.

To avoid the CPU having to wait for its next timeslot it is desirable to have a mechanism for ensuring that the CPU always gets the next available timeslot without incurring any latency on the non-CPU timeslots.

This can be done by defining each timeslot as consisting of a CPU access preceding a non-CPU access. Each timeslot will last 6 cycles i.e. a CPU access of 3 cycles and a non-CPU access of 3 cycles. This is exactly the interleaving behaviour outlined in Section 20.7.2.2. If the CPU does not require an access, the timeslot will take 3 or 4 and the timeslot rotation will go faster. A summary is given in Table 114.

TABLE 114 Timeslot access times. Access Duration Explanation CPU access + 3 + 3 = 6 Interleaved access non-CPU access cycles non-CPU access 4 cycles Access and preceding access both to shared read bus non-CPU access 3 cycles Access and preceding access not both to shared read bus CDU write access 3 + 2 + 2 + Page mode select signal 2 = 9 cycles is clocked at 160 MHz

CDU write accesses require 9 cycles. CDU write accesses preceded by a CPU access require 12 cycles. CDU timeslots therefore take longer than all other DIU requestors timeslots.

With a 256 cycle rotation there can be 42 accesses of 6 cycles.

For low scale factor applications, it is desirable to have more timeslots available in the same 256 5 cycle rotation. So two counters of 4-bits each are defined allowing the CPU to get a maximum of (CPUPreAccessTimeslots+1) pre-accesses for every (CPUTotalTimeslots+1) main slots. A timeslot counter starts at CPUTotalTimeslots and decrements every timeslot, while another counter starts at CPUPreAccessTimeslots and decrements every timeslot in which the CPU uses its access. When the CPU pre-access counter goes to zero before CPUTotalTimeslots, no further CPU accesses are allowed. When the CPUTotalTimeslots counter reaches zero both counters are reset to their respective initial values.

The CPU is not included in the list of SoPEC DIU requesters, Table, for the main timeslot allocations. The CPU cannot therefore be allocated main timeslots. It relies on pre-accesses in advance of such slots as the sole method for DRAM transfers.

CPU access to DRAM can never be fully disabled, since to do so would render SoPEC inoperable. Therefore the CPUPreAccessTimeslots and CPUTotalTimeslots register values are interpreted as follows: In each succeeding window of (CPUTotalTimeslots+1) slots, the maximum quota of CPU pre-accesses allowed is (CPUPreAccessTimeslots+1). The “+1” implementations mean that the CPU quota cannot be made zero.

The various modes of operation are summarised in Table 115 with a nominal rotation period of 256 cycles.

TABLE 115 CPU timeslot allocation modes with nominal rotation period of 256 cycles Nominal Timeslot Number of Access Type duration timeslots Notes CPU Pre-access 6 42 Each access is i.e. cycles timeslots CPU + non-CPU. CPUPreAccessTimeslots = If CPU does not CPUTotalTimeslots use a timeslot then rotation is faster. Fractional CPU 4 or 6 42-64 Each CPU + non- Pre-access i.e. cycles timeslots CPU access CPUPreAccessTimeslots < requires a 6 CPUTotalTimeslots cycle timeslot. Individual non- CPU timeslots take 4 cycles if current access and pre- ceding access are both to shared read bus. Individual non- CPU timeslots take 3 cycles if current ac- cess and pre- ceding access are not both to shared read bus. 20.10.4 CDU Accesses

As indicated in Section 20.10.3, CDU write accesses require 9 cycles. CDU write accesses preceded by a CPU access require 12 cycles. CDU timeslots therefore take longer than all other DIU requestors timeslots. This means that when a write timeslot is unused it cannot be re-allocated to a CDU write as CDU accesses take 9 cycles. The write accesses which the CDU write could otherwise replace require only 3 or 4 cycles.

Unused CDU write accesses can be replaced by any other write access according to 20.10.6.1 Unused write timeslots allocation on page 247.

20.10.5 Refresh Controller

Refresh is not included in the list of SoPEC DIU requesters, Table, for the main timeslot allocations. Timeslots cannot therefore be allocated to refresh.

The DRAM must be refreshed every 3.2 ms. Refresh occurs row at a time over 5120 rows of 2 parallel 10 Mbit instances. A refresh operation must therefore occur every 100 cycles. The refresh_period register has a default value of 99. Each refresh takes 3 cycles.

A refresh counter will count down the number of cycles between each refresh. When the down-counter reaches 0, the refresh controller will issue a refresh request and the down-counter is reloaded with the value in refresh_period and the count-down resumes immediately. Allocation of main slots must take into account that a refresh is required at least once every 100 cycles. Refresh is included in the unused read and write timeslot allocation. If unused timeslot allocation results in refresh occurring early by N cycles, then the refresh counter will have counted down to N. In this case, the refresh counter is reset to refresh_period and the count-down recommences. Refresh can be preceded by a CPU access in the same way as any other access. This is controlled by the CPUPreAccessTimeslots and CPUTotalTimeslots configuration registers. Refresh will therefore not affect CPU performance. A sequence of accesses including refresh might therefore be CPU, refresh, CPU, actual timeslot.

20.10.6 Allocating Unused Timeslots

Unused slots are re-allocated separately depending on whether the unused access was a read access or a write access. This is best-effort traffic. Only unused non-CPU accesses are re-allocated.

20.10.6.1 Unused Write Timeslots Allocation

Unused write timeslots are re-allocated according to a fixed priority order shown in Table 116.

TABLE 116 Unused write timeslot priority order Priority Name Order SCB(W) 1 SFU(W) 2 DWU 3 Unused read timeslot allocation 4

CDU write accesses cannot be included in the unused timeslot allocation for write as CDU accesses take 9 cycles. The write accesses which the CDU write could otherwise replace require only 3 or 4 cycles.

Unused write timeslot allocation occurs two timeslots in advance as noted in Section 20.10.2. If the units at priorities 1-3 are not requesting then the timeslot is re-allocated according to the unused read timeslot allocation scheme described in Section 20.10.6.2. However, the unused read timeslot allocation will occur when the current timeslot pointer of FIG. 96 reaches the timeslot i.e. it will not occur in advance.

20.10.6.2 Unused Read Timeslots Allocation

Unused read timeslots are re-allocated according to a two level round-robin scheme. The SoPEC Units included in read timeslot re-allocation is shown in Table 117.

TABLE 117 Unused read timeslot allocation Name SCB(R) CDU(R) CFU LBD SFU(R) TE(TD) TE(TFS) HCU DNC LLU PCU CPU/Refresh

Each SoPEC requestor has an associated bit, ReadRoundRobinLevel, which indicates whether it is in level 1 or level 2 round-robin.

TABLE 118 Read round-robin level selection Level Action ReadRoundRobinLevel = 0 Level 1 ReadRoundRobinLevel = 1 Level 2

A pointer points to the most recent winner on each of the round-robin levels. Re-allocation is carried out by traversing level 1 requesters, starting with the one immediately succeeding the last level 1 winner. If a requesting unit is found, then it wins arbitration and the level 1 pointer is shifted to its position. If no level 1 unit wants the slot, then level 2 is similarly examined and its pointer adjusted.

Since refresh occupies a (shared) position on one of the two levels and continually requests access, there will always be some round-robin winner for any unused slot.

20.10.6.2.1 Shared CPU/Refresh Round-Robin Position

Note that the CPU can conditionally be allowed to take part in the unused read round-robin scheme. Its participation is controlled via the configuration bit EnableCPURoundRobin. When this bit is set, the CPU and refresh share a joint position in the round-robin order, shown in Table. When cleared, the position is occupied by refresh alone.

If the shared position is next in line to be awarded an unused non-CPU read/write slot, then the CPU will have first option on the slot. Only if the CPU doesn't want the access, will it be granted to refresh. If the CPU is excluded from the round robin, then any awards to the position benefit refresh.

20.11 Guidelines for Programming the DIU

Some guidelines for programming the DIU arbitration scheme are given in this section together with an example.

20.11.1 Circuit Latency

Circuit latency is a fixed service delay which is incurred, as and from the acceptance by the DIU arbitration logic of a block's pending read/write request. It is due to the processing time of the request, readying the data, plus the DRAM access time. Latencies differ for read and write requests. See Tables 79 and 80 for respective breakdowns.

If a requesting block is currently stalled, then the longest time it will have to wait between issuing a new request for data and actually receiving it would be its timeslot period, plus the circuit latency overhead, along with any intervening non-standard slot durations, such as refresh and CDU(W). In any case, a stalled block will always incur this latency as an additional overhead, when coming out of a stall.

In the case where a block starts up or unstalls, it will start processing newly-received data at a time beyond its serviced timeslot equivalent to the circuit latency. If the block's timeslots are evenly spaced apart in time to match its processing rate, (in the hope of minimising stalls,) then the earliest that the block could restall, if not re-serviced by the DIU, would be the same latency delay beyond its next timeslot occurrence. Put another way, the latency incurred at start-up pushes the potential DIU-induced stall point out by the same fixed delta beyond each successive timeslot allocated to the block. This assumes that a block re-requests access well in advance of its upcoming timeslots. Thus, for a given stall-free run of operation, the circuit latency overhead is only incurred inititially when unstalling.

While a block can be stalled as a result of how quickly the DIU services its DRAM requests, it is also prone to stalls caused by its upstream or downstream neighbours being able to supply or consume data which is transferred between the blocks directly, (as opposed to via the DIU). Such neighbour-induced stalls, often occurring at events like end of line, will have the effect that a block's DIU read buffer will tend to fill, as the block stops processing read data. Its DIU write buffer will also tend to fill, unable to despatch to DRAM until the downstream block frees up shared-access DRAM locations. This scenario is beneficial, in that when a block unstalls as a result of its neighbour releasing it, then that block's read/write DIU buffers will have a fill state less likely to stall it a second time, as a result of DIU service delays.

A block's slots should be scheduled with a service guarantee in mind. This is dictated by the block's processing rate and hence, required access to the DRAM. The rate is expressed in terms of bits per cycle across a processing window, which is typically (though not always) 256 cycles. Slots should be evenly interspersed in this window (or “rotation”) so that the DIU can fulfill the block's service needs.

The following ground rules apply in calculating the distribution of slots for a given non-CPU block:—

-   -   The block can, at maximum, suffer a stall once in the rotation,         (i.e. unstall and restall) and hence incur the circuit latency         described above.

This rule is, by definition, always fulfilled by those blocks which have a service requirement of only

-   -   1 bit/cycle (equivalent to 1 slot/rotation) or fewer. It can be         shown that the rule is also satisfied by those blocks requiring         more than 1 bit/cycle. See Section 20.12.1 Slot Distributions         and Stall Calculations for Individual Blocks, on page 255.     -   Within the rotation, certain slots will be unavailable, due to         their being used for refresh. (See Section 20.11.2 Refresh         latencies)     -   In programming the rotation, account must be taken of the fact         that any CDU(W) accesses will consume an extra 6 cycles/access,         over and above the norm, in CPU pre-access mode, or 5         cycles/access without pre-access.     -   The total delay overhead due to latency, refreshes and CDU(W)         can be factored into the service guarantee for all blocks in the         rotation by deleting once, (i.e. reducing the rotation window,)         that number of slots which equates to the cumulative duration of         these various anomalies.     -   The use of lower scale factors will imply a more frequent demand         for slots by non-CPU blocks. The percentage of slots in the         overall rotation which can therefore be designated as CPU         pre-access ones should be calculated last, based on what can be         accommodated in the light of the non-CPU slot need.

Read latency is summarised below in Table 119.

TABLE 119 Read latency Non-CPU read access latency Duration non-CPU read requestor internally generates DIU 1 cycle request register the non- CPU read request 1 cycle complete the arbitration of the request 1 cycle transfer the read address to the DRAM 1 cycle DRAM read latency 1 cycle register the DRAM read data in DIU 1 cycle register the 1st 64-bits of read data in requester 1 cycle register the 2nd 64-bits of read data in requester 1 cycle register the 3rd 64-bits of read data in requester 1 cycle register the 4th 64-bits of read data in requester 1 cycle TOTAL 10 cycles

Write latency is summarised in Table 120.

TABLE 120 Write latency Non-CPU write access latency Duration non-CPU write requestor internally generates DIU request 1 cycle register the non-CPU write request 1 cycle complete the arbitration of the request 1 cycle transfer the acknowledge to the write requester 1 cycle transfer the 1st 64 bits of write data to the DIU 1 cycle transfer the 2nd 64 bits of write data to the DIU 1 cycle transfer the 3rd 64 bits of write data to the DIU 1 cycle transfer the 4th 64 bits of write data to the DIU 1 cycle Write to DRAM with locally registered write data 1 cycle TOTAL 9 cycles

Timeslots removed to allow for read latency will also cover write latency, since the former is the larger of the two.

20.11.2 Refresh Latencies

The number of allocated timeslots for each requester needs to take into account that a refresh must occur every 100 cycles. This can be achieved by deleting timeslots from the rotation since the number of timeslots is made programmable.

Refresh is preceded by a CPU access in the same way as any other access. This is controlled by the CPUPreAccessTimeslots and CPUTotalTimeslots configuration registers. Refresh will therefore not affect CPU performance.

As an example, in CPU pre-access mode each timeslot will last 6 cycles. If the timeslot rotation has 50 timeslots then the rotation will last 300 cycles. The refresh controller will trigger a refresh every 100 cycles. Up to 47 timeslots can be allocated to the rotation ignoring refresh. Three timeslots deleted from the 50 timeslot rotation will allow for the latency of a refresh every 100 cycles.

20.11.3 Ensuring Sufficient DNC and PCU Access

PCU command reads from DRAM are exceptional events and should complete in as short a time as possible. Similarly, we must ensure there is sufficient free bandwidth for DNC accesses e.g. when clusters of dead nozzles occur. In Table DNC is allocated 3 times average bandwidth. PCU and DNC can also be allocated to the level 1 round-robin allocation for unused timeslots so that unused timeslot bandwidth is preferentially available to them.

20.11.4 Basing Timeslot Allocation on Peak Bandwidths

Since the embedded DRAM provides sufficient bandwidth to use 1:1 compression rates for the CDU and LBD, it is possible to simplify the main timeslot allocation by basing the allocation on peak bandwidths. As combined bi-level and tag bandwidth at 1:1 scaling is only 5 bits/cycle, we will usually only consider the contone scale factor as the variable in determining timeslot allocations.

If slot allocation is based on peak bandwidth requirements then DRAM access will be guaranteed to all SoPEC requesters. If we do not allocate slots for peak bandwidth requirements then we can also allow for the peaks deterministically by adding some cycles to the print line time.

20.11.5 Adjacent Timeslot Restrictions

20.11.5.1 Non-CPU Write Adjacent Timeslot Restrictions

Non-CPU write requesters should not be assigned adjacent timeslots as described in Section 20.7.2.3. This is because adjacent timeslots assigned to non-CPU requestors would require two sets of 256-bit write buffers and multiplexors to connect two write requestors simultaneously to the DIU. Only one 256-bit write buffer and multiplexor is implemented. Recall from section 20.7.2.3 on page 238 that if adjacent non-CPU writes are attempted, that the second write of any such pair will be disregarded and re-allocated under the unused read scheme.

20.11.5.2 Same DIU Requestor Adjacent Timeslot Restrictions

All DIU requesters have state-machines which request and transfer the read or write data before requesting again. From FIG. 90 read requests have a minimum separation of 9 cycles. From FIG. 92 write requests have a minimum separation of 7 cycles. Therefore adjacent timeslots should not be assigned to a particular DIU requester because the requester will not be able to make use of all these slots.

In the case that a CPU access precedes a non-CPU access timeslots last 6 cycles so write and read requesters can only make use of every second timeslot. In the case that timeslots are not preceded by CPU accesses timeslots last 4 cycles so the same write requester can use every second timeslot but the same read requestor can use only every third timeslot. Some DIU requestors may introduce additional pipeline delays before they can request again. Therefore timeslots should be separated by more than the minimum to allow a margin.

20.11.6 Line Margin

The SFU must output 1 bit/cycle to the HCU. Since HCUNumDots may not be a multiple of 256 bits the last 256-bit DRAM word on the line can contain extra zeros. In this case, the SFU may not be able to provide 1 bit/cycle to the HCU. This could lead to a stall by the SFU. This stall could then propagate if the margins being used by the HCU are not sufficient to hide it. The maximum stall can be estimated by the calculation: DRAM service period—X scale factor * dots used from last DRAM read for HCU line.

Similarly, if the line length is not a multiple of 256-bits then e.g. the LLU could read data from DRAM which contains padded zeros. This could lead to a stall. This stall could then propagate if the page margins cannot hide it.

A single addition of 256 cycles to the line time will suffice for all DIU requesters to mask these stalls.

20.12 Example Outline DIU Programming

TABLE 121 Timeslot allocation based on peak bandwidth Peak Bandwidth which must be Block supplied MainTimeslots Name Direction (bits/cycle) allocated SCB R W 0.734⁷ 1 CDU R 0.9 (SF = 6), 1 (SF = 6) 2 (SF = 4) 2 (SF = 4) W 1.8 (SF = 6),⁸ 2 (SF = 6) 4 (SF = 4) 4 (SF = 4) CFU R 5.4 (SF = 6), 6 (SF = 6) 8 (SF = 4) 8 (SF = 4) LBD R 1 1 SFU R 2 2 W 1 1 TE(TD) R 1.02 1 TE(TFS) R 0.093 0 HCU R 0.074 0 DNC R 2.4 3 DWU W 6 6 LLU R 8 8 PCU R 1 1 TOTAL 33 (SF = 6) 38 (SF = 4) ⁷The SCB figure of 0.734 bits/cycl applies to multi-SoPEC systems. For single-SoPEC systems, the figure is 0.050 bits/cycle. ⁸Bandwidth for CDU(W) is peak value. Because of 1.5 buffering in DRAM, peak CDU(W) b/w equals 2 × average CDU(W) b/w. For CDU(R), peak b/w = average CDU(R) b/w.

Table 121 shows an allocation of main timeslots based on the peak bandwidths of Table The bandwidth required for each unit is calculated allowing extra cycles for read and write circuit latency for each access requiring a bandwidth of more than 1 bit/cycle. Fractional bandwidth is supplied via unused read slots.

The timeslot rotation is 256 cycles. Timeslots are deleted from the rotation to allow for circuit latencies for accesses of up to 1 bit per cycle i.e. 1 timeslot per rotation.

Example 1: Scale-Factor=6

Program the MainTimeslot configuration register (Table) for peak required bandwidths of SoPEC Units according to the scale factor.

Program the read round-robin allocation to share unused read slots. Allocate PCU, DNC, HCU and TFS to level 1 read round-robin.

-   Assume scale-factor of 6 and peak bandwidths from Table     -   Assign all DIU requestors except TE(TFS) and HCU to multiples of         1 timeslot, as indicated in Table, where each timeslot is 1         bit/cycle. This requires 33 timeslots.     -   No timeslots are explicitly allocated for the fractional         bandwidth requirements of TE(TFS) and HCU accesses. Instead,         these units are serviced via unused read slots.     -   Allow 3 timeslots to allow for 3 refreshes in the rotation.     -   Therefore, 36 scheduled slots are used in the rotation for main         timeslots and refreshes, some or all of which may be able to         have a CPU pre-access, provided they fit in the rotation window.     -   Each of the 2 CDU(W) accesses requires 9 cycles. Per access,         this implies an overhead of 1 slot (12 cycles instead of 6) in         pre-access mode, or 1.25 slots (9 cycles instead of 4) for no         pre-access. The cumulative overhead of the two accesses is         either 2 slots (pre-access) or 3 slots (no pre-access).     -   Assuming all blocks require a service guarantee of no more than         a single stall across 256 bits, allow 10 cycles for read         latency, which also takes care of 9-cycle write latency. This         can be accounted for by reserving 2 six-cycle slots (CPU         pre-access) or 3 four-cycle slots (no pre-access).     -   Assume a 256 cycle timeslot rotation.     -   CDU(W) and read latency reduce the number of available cycles in         a rotation to: 256−2×6−2×6=232 cycles (CPU pre-access) or         256−3×4−3×4=232 cycles (no pre-access).     -   As a result, 232 cycles available for 36 accesses implies each         access can take 232/36=6.44 cycles maximum. So, all accesses can         have a pre-access.     -   Therefore the CPU achieves a pre-access ratio of 36/36=100% of         slots in the rotation.         Example 2: Scale-Factor=4

Program the MainTimeslot configuration register (Table) for peak required bandwidths of SoPEC Units according to the scale factor. Program the read round-robin allocation to share unused read slots. Allocate PCU, DNC, HCU and TFS to level 1 read round-robin.

-   -   Assume scale-factor of 4 and peak bandwidths from Table     -   Assign all DIU requestors except TE(TFS) and HCU multiples of 1         timeslot, as indicated in Table, where each timeslot is 1         bit/cycle. This requires 38 timeslots.     -   No timeslots are explicitly allocated for the fractional         bandwidth requirements of TE(TFS) and HCU accesses. Instead,         these units are serviced via unused read slots.     -   Allow 3 timeslots to allow for 3 refreshes in the rotation.     -   Therefore, 41 scheduled slots are used in the rotation for main         timeslots and refreshes, some or all of which can have a CPU         pre-access, provided they fit in the rotation window.     -   Each of the 4 CDU(W) accesses requires 9 cycles. Per access,         this implies an overhead of 1 slot (12 cycles instead of 6) for         pre-access mode, or 1.25 slots (9 cycles instead of 4) for no         pre-access. The cumulative overhead of the four accesses is         either 4 slots (pre-access) or 5 slots (no pre-access).     -   Assuming all blocks require a service guarantee of no more than         a single stall across 256 bits, allow 10 cycles for read         latency, which also takes care of 9-cycle write latency. This         can be accounted for by reserving 2 six-cycle slots (CPU         pre-access) or 3 four-cycle slots (no pre-access).     -   Assume a 256 cycle timeslot rotation.     -   CDU(W) and read latency reduce the number of available cycles in         a rotation to: 256−4×6−2×6=220 cycles (CPU pre-access) or         256−5×4−3×4=224 cycles (no pre-access).     -   As a result, between 220 and 224 cycles are available for 41         accesses, which implies each access can take between 220/41=5.36         cycles and 224/41=5.46 cycles.     -   Work out how many slots can have a pre-access: For the lower         number of 220 cycles, this implies (41−n)*6+n*4<=220, where         n=number of slots with no pre-access cycle. Solving the equation         gives n>=13. Check answer: 28*6+0.13*4=220.     -   So 28 slots out of the 41 in the rotation can have CPU         pre-accesses.     -   The CPU thus achieves a pre-access ratio of 28/41=68.3% of slots         in the rotation.         20.12.1 Slot Distributions and Stall Calculations for Individual         Blocks

The following sections show how the slots for blocks with a service requirement greater than 1 bit/cycle should be distributed. Calculations are included to check that such blocks will not suffer more than one stall per rotation.

20.12.1.1 SFU

This has 2 bits/cycle on read but this is two separate channels of 1 bit/cycle sharing the same DIU interface so it is effectively 2 channels each of 1 bit/cycle so allowing the same margins as the LBD will work.

20.12.1.2 DWU

The DWU has 12 double buffers in each of the 6 colour planes, odd and even. These buffers are filled by the DNC and will request DIU access when double buffers fill. The DNC supplies 6 bits to the DWU every cycle (6 odd in one cycle, 6 even in the next cycle). So the service deadline is 512 cycles, given 6 accesses per 256-cycle rotation.

20.12.1.3 CFU

Here the requirement is that the DIU stall should be less than the time taken for the CFU to consume one third of its triple buffer. The total DIU stall=refresh latency+extra CDU(W) latency+read circuit latency=3+5 (for 4 cycle timeslots)+10=18 cycles. The CFU can consume its data at 8 bits/cycle at SF=4. Therefore 256 bits of data will last 32 cycles so the triple buffer is safe. In fact we only need an extra 144 bits of buffering or 3×64 bits. But it is safer to have the full extra 256 bits or 4×64 bits of buffering.

20.12.1.4 LLU

The LLU has 2 channels, each of which could request at 6 bits/106 MHz channel or 4 bits/160 MHz cycle, giving a total of 8 bits/160 MHz cycle. The service deadline for each channel is 256×106 MHz cycles, i.e. all 6 colours must be transferred in 256 cycles to feed the printhead. This equates to 384×160 MHz cycles.

Over a span of 384 cycles, there will be 6 CDU(W) accesses, 4 refreshes and one read latency encountered at most. Assuming CPU pre-accesses for these occurrences, this means the number of available cycles is given by 384−6×6−4×6−10=314 cycles.

For a CPU pre-access slot rate of 50%, 314 cycles implies 31 CPU and 63 non-CPU accesses (31×6+32×4=314). For 12 LLU accesses interspersed amongst these 63 non-CPU slots, implies an LLU allocation rate of approximately one slot in 5.

If the CPU pre-access is 100% across all slots, then 314 cycles gives 52 slots each to CPU and non-CPU accesses, (52×6=312 cycles). Twelve accesses spread over 52 slots, implies a 1-in-4 slot allocation to the LLU.

The same LLU slot allocation rate (1 slot in 5, or 1 in 4) can be applied to programming slots across a 256-cycle rotation window. The window size does not affect the occurrence of LLU slots, so the 384-cycle service requirement will be fulfilled.

20.12.1.5 DNC

This has a 2.4 bits/cycle bandwidth requirement. Each access will see the DIU stall of 18 cycles. 2.4 bits/cycle corresponds to an access every 106 cycles within a 256 cycle rotation. So to allow for DIU latency we need an access every 106-18 or 88 cycles. This is a bandwidth of 2.9 bits/cycle, requiring 3 timeslots in the rotation.

20.12.1.6 CDU

The JPEG decoder produces 8 bits/cycle. Peak CDUR[ead] bandwidth is 4 bits/cycle (SF=4), peak CDUW[rite] bandwidth is 4 bits/cycle (SF=4). both with 1.5 DRAM buffering.

The CDU(R) does a DIU read every 64 cycles at scale factor 4 with 1.5 DRAM buffering. The delay in being serviced by the DIU could be read circuit latency (10)+refresh (3)+extra CDU(W) cycles (6)=19 cycles. The JPEG decoder can consume each 256 bits of DIU-supplied data at 8 bits/cycle, i.e. in 32 cycles. If the DIU is 19 cycles late (due to latency) in supplying the read data then the JPEG decoder will have finished processing the read data 32+19=49 cycles after the DIU access. This is 64-49=15 cycles in advance of the next read. This 15 cycles is the upper limit on how much the DIU read service can further be delayed, without causing a stall. Given this margin, a stall on the read side will not occur.

On the write side, for scale factor 4, the access pattern is a DIU writes every 64 cycles with 1.5 DRAM buffereing. The JPEG decoder runs at 8 bits cycle and consumes 256 bits in 32 cycles. The CDU will not stall if the JPEG decode time (32)+DIU stall (19)<64, which is true.

20.13 CPU DRAM Access Performance

The CPU's share of the timeslots can be specified in terms of guaranteed bandwidth and average bandwidth allocations.

-   -   The CPU's access rate to memory depends on the CPU read access         latency i.e. the time between the CPU making a request to the         DIU and receiving the read data back from the DIU.     -   how often it can get access to DIU timeslots.

Table estimated the CPU read latency as 6 cycles.

How often the CPU can get access to DIU timeslots depends on the access type. This is summarised in Table 122.

TABLE 122 CPU DRAM access performance Nominal Access Timeslot CPU DRAM Type Duration access rate Notes CPU Pre- 6 Lower bound CPU can access every access cycles (guaranteed timeslot. bandwidth) is 160 MHz/ 6 = 26.27 MHz Frac- 4 or 6 Lower bound CPU accesses precede a tional cycles (guaranteed fraction N of timeslots CPU Pre- bandwidth) where N = C/T. access is (160 C = CPUPreAccessTimeslots MHz * N/P) T = CPUTotalTimeslots P = (6*C + 4*(T − C))/T

In both CPU Pre-access and Fractional CPU Pre-access modes, if the CPU is not requesting the timeslots will have a duration of 3 or 4 cycles depending on whether the current access and preceding access are both to the shared read bus. This will mean that the timeslot rotation will run faster and more bandwidth is available.

If the CPU runs out of its instruction cache then instruction fetch performance is only limited by the on-chip bus protocol. If data resides in the data cache then 160 MHz performance is achieved. Accessing memory mapped registers, PSS or ROM with a 3 cycle bus protocol (address cycle+data cycle) gives 53 MHz performance.

Due to the action of CPU caching, some bandwidth limiting of the CPU in Fractional CPU Pre-access mode is expected to have little or no impact on the overall CPU performance.

20.14 Implementation

The DRAM Interface Unit (DIU) is partitioned into 2 logical blocks to facilitate design and verification.

-   a. The DRAM Arbitration Unit (DAU) which interfaces with the SoPEC     DIU requesters. -   b. The DRAM Controller Unit (DCU) which accesses the embedded DRAM.

The basic principle in design of the DIU is to ensure that the eDRAM is accessed at its maximum rate while keeping the CPU read access latency as low as possible.

The DCU is designed to interface with single bank 20 Mbit IBM Cu-11 embedded DRAM performing random accesses every 3 cycles. Page mode burst of 4 write accesses, associated with the CDU, are also supported.

The DAU is designed to support interleaved accesses allowing the DRAM to be accessed every 3 cycles where back-to-back accesses do not occur over the shared 64-bit read data bus.

20.14.1 DIU Partition

20.14.2 Definition of DCU IO

TABLE 123 DCU interface Port Name Pins I/O Description Clocks and Resets pclk 1 In SoPEC Functional clock dau_dcu_reset_n 1 In Active-low, synchronous reset in pclk domain. Incorporates DAU hard and soft resets. Inputs from DAU dau_dcu_msn2stall 1 In Signal indicating from DAU Arbitration Logic which when asserted stalls DCU in MSN2 state. dau_dcu_adr[21:5] 17 In Signal indicating the address for the DRAM access. This is a 256- bit aligned DRAM address. dau_dcu_rwn 1 In Signal indicating the direction for the DRAM access (1 = read, 0 = write). dau_dcu_cduwpage 1 In Signal indicating if access is a CDU write page mode access (1 = CDU page mode, 0 = not CDU page mode). dau_dcu_refresh 1 In Signal indicating that a refresh command is to be issued. If asserted dau_dcu_adr, dau_dcu_rwn and dau_dcu_cduwpage are ignored. dau_dcu_wdata 256 In 256-bit write data to DCU dau_dcu_wmask 32 In Byte encoded write data mask for 256-bit dau_dcu_wdata to DCU Polarity: A “1” in a bit field of dau_dcu_wmask means that the corre- sponding byte in the 256- bit dau_dcu_wdata is written to DRAM. Outputs to DAU dcu_dau_adv 1 Out Signal indicating to DAU to supply next command to DCU dcu_dau_wadv 1 Out Signal indicating to DAU to ini tiate next non-CPU write dcu_dau_(—) 1 Out Signal indicating that refreshcomplete the DCU has completed a refresh. dcu_dau_rdata 256 Out 256-bit read data from DCU. dcu_dau_rvalid 1 Out Signal indicating valid read data on dcu_dau_rdata. 20.14.3 DRAM Access Types

The DRAM access types used in SoPEC are summarised in Table 124. For a refresh operation the DRAM generates the address internally.

TABLE 124 SoPEC DRAM access types Type Access Read Random 256-bit read Write Random 256-bit write with byte write masking Page mode write for burst of 4 256-bit words with byte write masking Refresh Single refresh 20.14.4 Constructing the 20 Mbit DRAM From Two 10 Mbit Instances

The 20 Mbit DRAM is constructed from two 10 Mbit instances. The address ranges of the two instances are shown in Table 125.

TABLE 125 Address ranges of the two 10 Mbit instances in the 20 Mbit DRAM Hex 256-bit Binary 256-bit Instance Address word address word address Instance0 First word 00000 0 0000 0000 0000 0000 in lower 10 Mbit Instance0 Last word 09FFF 0 1001 1111 1111 1111 in lower 10 Mbit Instance1 First word 0A000 0 1010 0000 0000 0000 in upper 10 Mbit Instance1 Last word 13FFF 1 0011 1111 1111 1111 in upper 10 Mbit

There are separate macro select signals, inst0_MSN and inst1_MSN, for each instance and separate dataout busses inst0_DO and inst1_DO, which are multiplexed in the DCU. Apart from these signals both instances share the DRAM output pins of the DCU.

The DRAM Arbitration Unit (DAU) generates a 17 bit address, dau_dcu_adr[21:5], sufficient to address all 256-bit words in the 20 Mbit DRAM. The upper 5 bits are used to select between the two memory instances by gating their MSN pins. If instance1 is selected then the lower 16-bits are translated to map into the 10 Mbit range of that instance. The multiplexing and address translation rules are shown in Table 126.

In the case that the DAU issues a refresh, indicated by dau_dcu_refresh, then both macros are selected. The other control signals

TABLE 126 Instance selection and address translation DAU Address bits Instance dau_dcu_refresh dau_dcu_adr[21:17] selected inst0_MSN inst1_MSN Address translation 0  <01010 Instance0 MSN 1 A[15:0] = dau_dcu_adr[20:5] >=01010 Instance1 1 MSN A[15:0] = dau_dcu_adr[21:5] − hA000 1 — Instance0 MSN MSN — and Instance1 dau_dcu_adr[21:5], dau_dcu_rwn and dau_dcu_cduwpage are ignored.

The instance selection and address translation logic is shown in FIG. 102.

The address translation and instance decode logic also increments the address presented to the DRAM in the case of a page mode write. Pseudo code is given below.

if rising_(—edge(dau)_dcu_valid) then //capture the address from the DAU next_cmdadr[21:5] = dau_dcu_adr[21:5] elsif pagemode_adr_inc == 1 then //increment the address next_cmdadr[21:5] = cmdadr[21:5] + 1 else next_cmdadr[21:5] = cmdadr[21:5] if rising_edge(dau_dcu_valid) then //capture the address from the DAU adr_var[21:5] := dau_dcu_adr[21:5] else adr_var[21:5] := cmdadr[21:5] if adr_var[21:17] < 01010 then //choose instance0 instance_sel = 0 A[15:0] = adr_var[20:5] else //choose instance1 instance_sel = 1 A[15:0] = adr_var[21:5] − hA000

Pseudo code for the select logic, SEL0, for DRAM Instance0 is given below.

//instance0 selected or refresh if instance_sel == 0 OR dau_dcu_refresh == 1 then inst0_MSN = MSN else inst0_MSN = 1

Pseudo code for the select logic, SEL1, for DRAM instance1 is given below.

//instance1 selected or refresh if instance_sel == 1 OR dau_dcu_refresh == 1 then inst1_MSN = MSN else inst1_MSN = 1

During a random read, the read data is returned, on dcu_dau_rdata, after time T_(acc), the random access time, which varies between 3 and 8 ns (see Table). To avoid any metastability issues the read data must be captured by a flip-flop which is enabled 2 pclk cycles or 12.5 ns after the DRAM access has been started. The DCU generates the enable signal dcu_dau_rvalid to capture dcu_dau_rdata.

The byte write mask dau_dcu_wmask[31:0] must be expanded to the bit write mask bitwritemask[255:0]needed by the DRAM.

20.14.5 DAU-DCU Interface Description

The DCU asserts dcu_dau_adv in the MSN2 state to indicate to the DAU to supply the next command. dcu_dau_adv causes the DAU to perform arbitration in the MSN2 cycle. The resulting command is available to the DCU in the following cycle, the RST state. The timing is shown in FIG. 103. The command to the DRAM must be valid in the RST and MSN1 states, or at least meet the hold time requirement to the MSN falling edge at the start of the MSN1 state.

Note that the DAU issues a valid arbitration result following every dcu_dau_adv pulse. If no unit is requesting DRAM access, then a fall-back refresh request will be issued. When dau_dcu_refresh is asserted the operation is a refresh and dau_dcu_adr, dau_dcu_rwn and dau_dcu_cduwpage are ignored.

The DCU generates a second signal, dcu_dau_wadv, which is asserted in the RST state.

This indicates to the DAU that it can perform arbitration in advance for non-CPU writes.

The reason for performing arbitration in advance for non-CPU writes is explained in “Command Multiplexor Sub-block

TABLE 136 Command Multiplexor Sub-block IO Definition Port name Pins I/O Description Clocks and Resets pclk 1 In System Clock prst_n 1 In System reset, synchro- nous active low DIU Read Interface to SoPEC Units <unit>_diu_radr[21:5] 17 In Read address to DIU 17 bits wide (256-bit aligned word). diu_<unit>_rack 1 Out Acknowledge from DIU that read request has been accepted and new read address can be placed on <unit>_diu_radr DIU Write Interface to SoPEC Units <unit>_diu_wadr[21:5] 17 In Write address to DIU except CPU, SCB, CDU 17 bits wide (256-bit aligned word) cpu_diu_wadr[21:4]] 22 In CPU Write address to DIU (128-bit aligned address.) cpu_diu_wmask 16 In Byte enables for CPU write. cdu_diu_wadr[21:3] 19 In CDU Write address to DIU 19 bits wide (64-bit aligned word) Addresses cannot cross a 256-bit word DRAM boundary. diu_<unit>_wack 1 Out Acknowledge from DIU that write request has been accepted and new write address can be placed on <unit>_diu_wadr Outputs to CPU Interface and Arbitration Logic sub-block re_arbitrate 1 Out Signalling telling the arbitration logic to choose the next arbitra- tion winner. re_arbitrate_wadv 1 Out Signal telling the arbitration logic to choose the next arbitra- tion winner for non- CPU writes 2 timeslots in advance Debug Outputs to CPU Configuration and Arbitration Logic Sub-block write_sel 5 Out Signal indicating the SoPEC Unit for which the current write trans- action is occurring. Encoding is described in Table . write_complete 1 Out Signal indicating that write transaction to SoPEC Unit indicated by write_sel is complete. Inputs from CPU Interface and Arbitration Logic sub-block arb_gnt 1 In Signal lasting 1 cycle which indicates arbi- tration has occurred and arb_sel is valid. arb_sel 5 In Signal indicating which requesting SoPEC Unit has won arbitration. Encoding is described in Table . dir_sel 2 In Signal indicating which sense of access asso- ciated with arb_sel 00: issue non-CPU write 01: read winner 10: write winner 11: refresh winner Inputs from Read Write Multiplexor Sub-block write_data_valid 2 In Signal indicating that valid write data is available for the current command. 00 = not valid 01 = CPU write data valid 10 = non-CPU write data valid 11 = both CPU and non- CPU write data valid wdata 256 In 256-bit non-CPU write data cpu_wdata 32 In 32-bit CPU write data Outputs to Read Write Multiplexor Sub-block write_data_accept 2 Out Signal indicating the Command Multiplexor has accepted the write data from the write multiplexor 00 = not valid 01 = accepts CPU write data 10 = accepts non-CPU write data 11 = not valid Inputs from DCU dcu_dau_adv 1 In Signal indicating to DAU to supply next command to DCU dcu_dau_wadv 1 In Signal indicating to DAU to initiate next non-CPU write Outputs to DCU dau_dcu_adr[21:5] 17 Out Signal indicating the address for the DRAM access. This is a 256- bit aligned DRAM address. dau_dcu_rwn 1 Out Signal indicating the direction for the DRAM access (1 = read, 0 = write). dau_dcu_cduwpage 1 Out Signal indicating if access is a CDU write page mode access (1 = CDU page mode, 0 = not CDU page mode). dau_dcu_refresh 1 Out Signal indicating that a refresh command is to be issued. If asserted dau_dcu_adr, dau_dcu_rwn and dau_dcu_cduwpage are ignored. dau_dcu_wdata 256 Out 256-bit write data to DCU dau_dcu_wmask 32 Out Byte encoded write data mask for 256- bit dau_dcu_wdata to DCU

The DCU state-machine can stall in the MSN2 state when the signal dau_dcu_msn2stall is asserted by the DAU Arbitration Logic,

The states of the DCU state-machine are summarised in Table 127.

TABLE 127 States of the DCU state-machine State Description RST Restore state MSN1 Macro select state 1 MSN2 Macro select state 2 20.14.6 DCU State Machines

The IBM DRAM has a simple SRAM like interface. The DRAM is accessed as a single bank. The state machine to access the DRAM is shown in FIG. 104.

The signal pagemode_adr_inc is exported from the DCU as dcu_dau_cduwaccept. dcu_dau_cduwaccept tells the DAU to supply the next write data to the DRAM

20.14.7 CU-11DRAM Timing Diagrams

The IBM Cu-11 embedded DRAM datasheet is referenced as [16].

Table 128 shows the timing parameters which must be obeyed for the IBM embedded DRAM.

TABLE 128 1.5 V Cu-11 DRAM a.c. parameters Symbol Parameter Min Max Units T_(set) Input setup to MSN/PGN 1 — ns T_(hld) Input hold to MSN/PGN 2 — ns T_(acc) Random access time 3 8 ns T_(act) MSN active time 8 100 k ns T_(res) MSN restore time 4 — ns T_(cyc) Random R/W cycle time 12 — ns T_(rfc) Refresh cycle time 12 — ns T_(accp) Page mode access time 1  3.9 ns T_(pa) PGN active time 1.6 — ns T_(pr) PGN restore time 1.6 — ns T_(pcyc) PGN cycle time 4 — ns T_(mprd) MSN to PGN restore delay 6 — ns T_(actp) MSN active for page mode 12 — ns T_(ref) Refresh period —  3.2 ms T_(pamr) Page active to MSN restore 4 — ns

The IBM DRAM is asynchronous. In SoPEC it interfaces to signals clocked on pclk. The following timing diagrams show how the timing parameters in Table 129 are satisfied in SoPEC.

20.14.8 Definition of DAU IO

TABLE 129 DAU interface Port Name Pins I/O Description Clocks and Resets pclk 1 In SoPEC Functional clock prst_n 1 In Active-low, synchronous reset in pclk domain dau_dcu_reset_n 1 Out Active-low, synchronous reset in pclk domain. This reset signal, exported to the DCU, incorporates the locally captured DAU version of hard reset (prst_n) and the soft reset configuration register bit “Reset”. CPU Interface cpu_adr 22 In CPU address bus for both DRAM and configuration register access. 9 bits (bits 10:2) are required to decode the configuration register address space. 22 bits can address the DRAM at byte level. DRAM addresses cannot cross a 256-bit word DRAM boundary. cpu_dataout 32 In Shared write data bus from the CPU for DRAM and configuration data diu_cpu_data 32 Out Configuration, status and debug read data bus to the CPU diu_cpu_(—) 1 Out Signal indicating the data debug_valid on the diu_cpu_data bus is valid debug data. cpu_rwn 1 In Common read/not-write signal from the CPU cpu_acode 2 In CPU access code signals. cpu_acode[0] - Program (0)/ Data (1) access cpu_acode[1] - User (0)/ Supervisor (1) access The DAU will only allow super- visor mode accesses to data space. cpu_diu_sel 1 In Block select from the CPU. When cpu_diu_sel is high both cpu_adr and cpu_dataout are valid diu_cpu_rdy 1 Out Ready signal to the CPU. When diu_cpu_rdy is high it indicates the last cycle of the access. For a write cycle this means cpu_dataout has been registered by the block and for a read cycle this means the data on diu_cpu_data is valid. diu_cpu_berr 1 Out Bus error signal to the CPU indicating an invalid access. DIU Read Interface to SoPEC Units <unit>_diu_rreq 1 In SoPEC unit requests DRAM read. A read request must be accompanied by a valid read address. <unit>_diu_(—) 17 In Read address to DIU radr[21:5] 17 bits wide (256-bit aligned word). Note: “<unit>” refers to non-CPU requesters only. CPU addresses are provided via “cpu_adr”. diu_<unit>_(—) 1 Out Acknowledge from DIU that rack read request has been accepted and new read address can be placed on <unit>_diu_radr diu_data 64 Out Data from DIU to SoPEC Units except CPU. First 64-bits is bits 63:0 of 256 bit word Second 64-bits is bits 127:64 of 256 bit word Third 64-bits is bits 191:128 of 256 bit word Fourth 64-bits is bits 255:192 of 256 bit word dram_cpu_data 256 Out 256-bit data from DRAM to CPU. diu_<unit>_(—) 1 Out Signal from DIU telling SoPEC rvalid Unit that valid read data is on the diu_data bus DIU Write Interface to SoPEC Units <unit>_diu_(—) 1 In SoPEC unit requests DRAM write. wreq A write request must be accompanied by a valid write address. Note: “<unit>” refers to non-CPU requesters only. <unit>_diu_(—) 17 In Write address to DIU except wadr[21:5] CPU, CDU 17 bits wide (256-bit aligned word) Note: “<unit>” refers to non-CPU requesters, excluding the CDU. scb_diu_(—) 8 In Byte write enables applicable wmask[7:0] to a given 64-bit quarter-word transferred from the SCB. Note that different mask values are used with each quarter-word. Requirement for the USB host core. diu_cpu_(—) 1 Out Flag indicating that the CPU write_rdy posted write buffer is empty. cpu_diu_(—) 1 In Write enable for the CPU wdatavalid posted write buffer. Also confirms that the CPU write data, address and mask are valid. cpu_diu_wdata 128 In CPU write data which is loaded into the posted write buffer. cpu_diu_(—) 18 In 128-bit aligned CPU write wadr[21:4] address. cpu_diu_(—) 16 In Byte enables for 128-bit CPU wmask[15:0] posted write. cdu_diu_(—) 19 In CDU Write address to DIU wadr[21:3] 19 bits wide (64-bit aligned word) Addresses cannot cross a 256-bit word DRAM boundary. diu_<unit>_(—) 1 Out Acknowledge from DIU that wack write request has been accepted and new write address can be placed on <unit>_diu_wadr <unit>_diu_(—) 64 In Data from SoPEC Unit to DIU data[63:0] except CPU. First 64-bits is bits 63:0 of 256 bit word Second 64-bits is bits 127:64 of 256 bit word Third 64-bits is bits 191:128 of 256 bit word Fourth 64-bits is bits 255:192 of 256 bit word Note: “<unit>” refers to non-CPU requesters only. <unit>_diu_(—) 1 In Signal from SoPEC Unit wvalid indicating that data on <unit>_diu_data is valid. Note: “<unit>” refers to non-CPU requesters only. Outputs to DCU dau_dcu_(—) 1 Out Signal indicating from DAU msn2stall Arbitration Logic which when de-asserted stalls DCU in MSN2 state. dau_dcu_(—) 17 Out Signal indicating the address adr[21:5] for the DRAM access. This is a 256-bit aligned DRAM address. dau_dcu_rwn 1 Out Signal indicating the direction for the DRAM access (1 = read, 0 = write). dau_dcu_(—) 1 Out Signal indicating if access cduwpage is a CDU write page mode access (1 = CDU page mode, 0 = not CDU page mode). dau_dcu_refresh 1 Out Signal indicating that a refresh command is to be issued. If asserted dau_dcu_cmd_adr, dau_dcu_rwn and dau_dcu_cduwpage are ignored. dau_dcu_wdata 256 Out 256-bit write data to DCU dau_dcu_wmask 32 Out Byte-encoded write data mask for 256-bit dau_dcu_wdata to DCU Polarity: A “1” in a bit field of dau_dcu_wmask means that the corresponding byte in the 256-bit dau_dcu_wdata is written to DRAM. Inputs from DCU dcu_dau_adv 1 In Signal indicating to DAU to supply next command to DCU dcu_dau_wadv 1 In Signal indicating to DAU to initiate next non-CPU write dcu_dau_(—) 1 In Signal indicating that the DCU refreshcomplete has completed a refresh. dcu_dau_rdata 256 In 256-bit read data from DCU. dcu_dau_rvalid 1 In Signal indicating valid read data on dcu_dau_rdata.

The CPU subsystem bus interface is described in more detail in Section 11.4.3. The DAU block will only allow supervisor-mode accesses to update its configuration registers (i.e. cpu_acode[1:0]=b11). All other accesses will result in diu_cpu_berr being asserted.

20.14.9 DAU Configuration Registers

TABLE 130 DAU configuration registers Address, # (DIU_base+) Register bits Reset Description Reset 0x00 Reset 1 0x1 A write to this register causes a reset of the DIU. This register can be read to indicate the reset state: 0 - reset in progress 1 - reset not in progress Refresh 0x04 RefreshPeriod 9 0x063 Refresh controller. When set to 0 refresh is off, other- wise the value indicates the number of cycles, less one, between each refresh. [Note that for a system clock frequency of 160 MHz, a value exceeding 0x63 (indicating a 100-cycle refresh period) should not be programmed, or the DRAM will malfunction.] Timeslot allocation and control 0x08 NumMainTimeslots 6 0x01 Number of main timeslots (1-64) less one 0x0C CPUPreAccessTimeslots 4 0x0 (CPUPreAccessTimeslots + 1) main slots out of a total of (CPUTotalTimeslots + 1) are preceded by a CPU access. 0x10 CPUTotalTimeslots 4 0x0 (CPUPreAccessTimeslots + 1) main slots out of a total of (CPUTotalTimeslots + 1) are preceded by a CPU access. 0x100-0x1FC MainTimeslot[63:0] 64 × 4 [63:1][3:0] = 0x0 Programmable main timeslots (up to [0][3:0] = 0xE 64 main timeslots). 0x200 ReadRoundRobinLevel 12 0x000 For each read requester plus refresh 0 = level1 of round-robin 1 = level2 of round-robin The bit order is defined in Table . 0x204 EnableCPURoundRobin 1 0x1 Allows the CPU to particpate in the unused read round-robin scheme. If disabled, the shared CPU/refresh round-robin position is dedicated solely to refresh. 0x208 RotationSync 1 0x1 Writing 0, followed by 1 to this bit allows the timeslot rotation to advance on a cycle basis which can be determined by the CPU. 0x20C minNonCPUReadAdr 12 0x800 12 MSBs of lowest DRAM address which may be read by non-CPU requesters. 0x210 minDWUWriteAdr 12 0x800 12 MSBs of lowest DRAM address which may be written to by the DWU. 0x214 minNonCPUWriteAdr 12 0x800 12 MSBs of lowest DRAM address which may be written to by non-CPU requesters other than the DWU. Debug 0x300 DebugSelect[11:2] 10 0x304 Debug address select. Indicates the address of the register to report on the diu_cpu_data bus when it is not otherwise being used. When this signal carries debug informa- tion the signal diu_cpu_debug_valid will be asserted. Debug: arbitration and performance 0x304 ArbitrationHistory 22 — Bit 0 = arb_gnt Bit 1 = arb_executed Bit 6:2 = arb_sel[4:0] Bit 12:7 = timeslot_number[5:0] Bit 15:13 = access_type[2:0] Bit 16 = back2back_non_cpu_write Bit 17 = sticky_back2back_non_cpu_write (Sticky version of same, cleared on reset.) Bit 18 = rotation_sync Bit 20:19 = rotation_state Bit 21 = sticky_invalid_non_cpu_adr See Section 20.14.9.2 DIU Debug for a description of the fields. Read only register. 0x308 DIUPerformance 31 — Bit 0 = cpu_diu_rreq Bit 1 = scb_diu_rreq Bit 2 = cdu_diu_rreq Bit 3 = cfu_diu_rreq Bit 4 = lbd_diu_rreq Bit 5 = sfu_diu_rreq Bit 6 = td_diu_rreq Bit 7 = tfs_diu_rreq Bit 8 = hcu_diu_rreq Bit 9 = dnc_diu_rreq Bit 10 = llu_diu_rreq Bit 11 = pcu_diu_rreq Bit 12 = cpu_diu_wreq Bit 13 = scb_diu_wreq Bit 14 = cdu_diu_wreq Bit 15 = sfu_diu_wreq Bit 16 = dwu_diu_wreq Bit 17 = refresh_req Bit 22:18 = read_sel[4:0] Bit 23 = read_complete Bit 28:24 = write_sel[4:0] Bit 29 = write_complete Bit 30 = dcu_dau_refreshcomplete See Section 20.14.9.2 DIU Debug for a description of the fields. Read only register. Debug DIU read requesters interface signals 0x30C CPUReadInterface 25 — Bit 0 = cpu_diu_rreq Bit 22:1 = cpu_adr[21:0] Bit 23 = diu_cpu_rack Bit 24 = diu_cpu_rvalid Read only register. 0x310 SCBReadInterface 20 Bit 0 = scb_diu_rreq Bit 17:1 = scb_diu_radr[21:5] Bit 18 = diu_scb_rack Bit 19 = diu_scb_rvalid Read only register. 0x314 CDUReadInterface 20 — Bit 0 = cdu_diu_rreq Bit 17:1 = cdu_diu_radr[21:5] Bit 18 = diu_cdu_rack Bit 19 = diu_cdu_rvalid Read only register. 0x318 CFUReadInterface 20 — Bit 0 = cfu_diu_rreq Bit 17:1 = cfu_diu_radr[21:5] Bit 18 = diu_cfu_rack Bit 19 = diu_cfu_rvalid Read only register. 0x31C LBDReadInterface 20 — Bit 0 = lbd_diu_rreq Bit 17:1 = lbd_diu_radr[21:5] Bit 18 = diu_lbd_rack Bit 19 = diu_lbd_rvalid Read only register. 0x320 SFUReadInterface 20 — Bit 0 = sfu_diu_rreq Bit 17:1 = sfu_diu_radr[21:5] Bit 18 = diu_sfu_rack Bit 19 = diu_sfu_rvalid Read only register. 0x324 TDReadInterface 20 — Bit 0 = td_diu_rreq Bit 17:1 = td_diu_radr[21:5] Bit 18 = diu_td_rack Bit 19 = diu_td_rvalid Read only register. 0x328 TFSReadInterface 20 — Bit 0 = tfs_diu_rreq Bit 17:1 = tfs_diu_radr[21:5] Bit 18 = diu_tfs_rack Bit 19 = diu_tfs_rvalid Read only register. 0x32C HCUReadInterface 20 — Bit 0 = hcu_diu_rreq Bit 17:1 = hcu_diu_radr[21:5] Bit 18 = diu_hcu_rack Bit 19 = diu_hcu_rvalid Read only register. 0x330 DNCReadInterface 20 — Bit 0 = dnc_diu_rreq Bit 17:1 = dnc_diu_radr[21:5] Bit 18 = diu_dnc_rack Bit 19 = diu_dnc_rvalid Read only register. 0x334 LLUReadInterface 20 — Bit 0 = llu_diu_rreq Bit 17:1 = lluu_diu_radr[21:5] Bit 18 = diu_llu_rack Bit 19 = diu_llu_rvalid Read only register. 0x338 PCUReadInterface 20 — Bit 0 = pcu_diu_rreq Bit 17:1 = pcu_diu_radr[21:5] Bit 18 = diu_pcu_rack Bit 19 = diu_pcu_rvalid Read only register. Debug DIU write requesters interface signals 0x33C CPUWriteInterface 27 — Bit 0 = cpu_diu_wreq Bit 22:1 = cpu_adr[21:0] Bit 24:23 = cpu_diu_wmask[1:0] Bit 25 = diu_cpu_wack Bit 26 = cpu_diu_wvalid Read only register. 0x340 SCBWriteInterface 20 — Bit 0 = scb_diu_wreq Bit 17:1 = scb_diu_wadr[21:5] Bit 18 = diu_scb_wack Bit 19 = scb_diu_wvalid Read only register. 0x344 CDUWriteInterface 22 — Bit 0 = cdu_diu_wreq Bit 19:1 = cdu_diu_wadr[21:3] Bit 20 = diu_cdu_wack Bit 21 = cdu_diu_wvalid Read only register. 0x348 SFUWriteInterface 20 — Bit 0 = sfu_diu_wreq Bit 17:1 = sfu_diu_wadr[21:5] Bit 18 = diu_sfu_wack Bit 19 = sfu_diu_wvalid Read only register. 0x34C DWUWriteInterface 20 — Bit 0 = dwu_diu_wreq Bit 17:1 = dwu_diu_wadr[21:5] Bit 18 = diu_dwu_wack Bit 19 = dwu_diu_wvalid Read only register. Debug DAU-DCU interface signals 0x350 DAU-DCUInterface 25 — Bit 16:0 = dau_dcu_adr[21:5] Bit 17 = dau_dcu_rwn Bit 18 = dau_dcu_cduwpage Bit 19 = dau_dcu_refresh Bit 20 = dau_dcu_msn2stall Bit 21 = dcu_dau_adv Bit 22 = dcu_dau_wadv Bit 23 = dcu_dau_refreshcomplete Bit 24 = dcu_dau_rvalid Read only register.

Each main timeslot can be assigned a SoPEC DIU requestor according to Table 131.

TABLE 131 SoPEC DIU requester encoding for main timeslots. Index Index Name (binary) (HEX) Write SCB(W) b0_0000 0x00 CDU(W) b0001 0x1 SFU(W) b0010 0x2 DWU b0011 0x3 Read SCB(R) b0100 0x4 CDU(R) b0101 0x5 CFU b0110 0x6 LBD b0111 0x7 SFU(R) b1000 0x8 TE(TD) b1001 0x9 TE(TFS) b1010 0xA HCU b1011 0xB DNC b1100 0xC LLU b1101 0xD PCU b1110 0xE

ReadRoundRobinLevel and ReadRoundRobinEnable registers are encoded in the bit order defined in Table 132.

TABLE 132 Read round-robin registers bit order Bit Name index SCB(R) 0 CDU(R) 1 CFU 2 LBD 3 SFU(R) 4 TE(TD) 5 TE(TFS) 6 HCU 7 DNC 8 LLU 9 PCU 10 CPU/Refresh 11 20.14.9.1 Configuration Register Reset State

The RefreshPeriod configuration register has a reset value of 0x063 which ensures that a refresh will occur every 100 cycles and the contents of the DRAM will remain valid.

The CPUPreAccessTimeslots and CPUTotalTimeslots configuration registers both have a reset value of 0x0. Matching values in these two registers means that every slot has a CPU pre-acess. NumMainTimeslots is reset to 0x1, so there are just 2 main timeslots in the rotation initially. These slots alternate between SCB writes and PCU reads, as defined by the reset value of MainTimeslot[63:0], thus respecting at reset time the general rule that adjacent non-CPU writes are not permitted.

The first access issued by the DIU after reset will be a refresh.

20.14.9.2 DIU Debug

External visibility of the DIU must be provided for debug purposes. To facilitate this debug registers are added to the DIU address space.

The DIU CPU system data bus diu_cpu_data[31:0] returns configuration and status register information to the CPU. When a configuration or status register is not being read by the CPU debug data is returned on diu_cpu_data[31:0] instead. An accompanying active high diu_cpu_debug_valid signal is used to indicate when the data bus contains valid debug data. The DIU features a DebugSelect register that controls a local multiplexor to determine which register is output on diu_cpu_data[31:0].

Three kinds of debug information are gathered:

-   a. The order and access type of DIU requesters winning arbitration.

This information can be obtained by observing the signals in the ArbitrationHistory debug register at DIU_Base+0x304 described in Table 133.

TABLE 133 ArbitrationHistory debug register description, DIU_base+0x304 Field name Bits Description arb_gnt 1 Signal lasting 1 cycle which is asserted in the cycle following a main arbitration or pre-arbitration. arb_executed 1 Signal lasting 1 cycle which indi- cates that an arbitration result has actually been executed. Is used to differentiate between *pre*- arbitration and *main* arbitration, both of which cause arb_gnt to be asserted. If arb_executed and arb_gnt are both high, then a main (executed) arbitration is indicated. arb_sel 5 Signal indicating which requesting SoPEC Unit has won arbitration. Encoding is described in Table . Refresh winning arbitration is indicated by access_type. timeslot_number 6 Signal indicating which main time- slot is either currently being serviced, or about to be serviced. The latter case applies where a main slot is preempted by a CPU pre-access or a scheduled refresh. access_type 3 Signal indicating the origin of the winning arbitration 000 = Standard CPU pre-access. 001 = Scheduled refresh. 010 = Standard non-CPU timeslot. 011 = CPU access via unused read/ write slot, re-allocated by round robin. 100 = Non-CPU write via unused write slot, re-allocated at prearbitration. 101 = Non-CPU read via unused read/ write slot, re-allocated by round robin. 110 = Refresh via unused read/write slot, re-allocated by round robin. 111 = CPU/Refresh access due to RotationSync = 0. back2back_(—) 1 Instantaneous indicator of attempted non_cpu_write illegal back-to-back non-CPU write. (Recall from section 20.7.2.3 on page 212 that the second write of any such pair is disregarded and re-allocated via the unused read round-robin scheme.) sticky_back2back_(—) 1 Sticky version of same, cleared on non_cpu_write reset. rotation_sync 1 Current value of the RotationSync configuration bit. rotation_state 2 These bits indicate the current status of pre-arbitation and main timeslot rotation, as a result of the RotationSync setting. 00 = Pre-arb enabled, rotation enabled. 01 = Pre-arb disabled, rotation enabled. 10 = Pre-arb disabled, rotation disabled. 11 = Pre-arb enabled, rotation disabled. 00 is the normal functional setting when RotationSync is 1. 01 indicates that pre-arbitration has halted at the end of its rotation because of RotationSync having been cleared. However the main arbitration has yet to finish its current rota- tion. 10 indicates that both pre-arb and the main rotation have halted, due to RotationSync being 0 and that only CPU accesses and refreshes are allowed. 11 indicates that RotationSync has just been changed from 0 to 1 and that pre-arbitration is being given a head start to look ahead for non- CPU writes, in advance of the main rotation starting up again. sticky_invalid_(—) 1 Sticky bit to indicate an attempted non_cpu_adr non-CPU access with an invalid ad- dress. Cleared by reset or by an explicit write by the CPU.

TABLE 134 arb_sel, read_sel and write_sel encoding Index Index Name (binary) (HEX) Write SCB(W) b0_0000 0x00 CDU(W) b0_0001 0x01 SFU(W) b0_0010 0x02 DWU b0_0011 0x03 Read SCB(R) b0_0100 0x04 CDU(R) b0_0101 0x05 CFU b0_0110 0x06 LBD b0_0111 0x07 SFU(R) b0_1000 0x08 TE(TD) b0_1001 0x09 TE(TFS) b0_1010 0x0A HCU b0_1011 0x0B DNC b0_1100 0x0C LLU b0_1101 0x0D PCU b0_1110 0x0E Refresh Refresh b0_1111 0x0F CPU CPU(R) b1_0000 0x10 CPU(W) b1_0001 0x11

The encoding for arb_sel is described in Table 134.

-   b. The time between a DIU requester requesting an access and     completing the access.

This information can be obtained by observing the signals in the DIUPerformance debug register at DIU_Base+0x308 described in Table 135. The encoding for read_sel and write_sel is described in Table. The data collected from DIUPerformance can be post-processed to count the number of cycles between a unit requesting DIU access and the access being completed.

TABLE 135 DIUPerformance debug register description, DIU_base+0x308 Field name Bits Description <unit>_diu_(—) 12 Signal indicating that SoPEC unit rreq requests DRAM read. <unit>_diu_(—) 5 Signal indicating that SoPEC unit wreq requests DRAM write. refresh_req 1 Signal indicating that refresh has requested a DIU access. read_sel[4:0] 5 Signal indicating the SoPEC Unit for which the current read transaction is occurring. Encoding is described in Table . read_complete 1 Signal indicating that read trans- action to SoPEC Unit indicated by read_sel is complete i.e. that the last read data has been output by the DIU. write_sel[4:0] 5 Signal indicating the SoPEC Unit for which the current write transaction is occurring. Encoding is described in Table . write_complete 1 Signal indicating that write trans- action to SoPEC Unit indicated by write_sel is complete i.e. that the last write data has been transferred to the DIU. dcu_refresh_(—) 1 Signal indicating that refresh has complete completed.

-   c. Interface Signals to DIU Requestors and DAU-DCU Interface

All interface signals with the exception of data busses at the interfaces between the DAU and DCU and DIU write and read requestors can be monitored in debug mode by observing debug registers DIU_Base+0x314 to DIU_Base+0x354.

20.14.10 DRAM Arbitration Unit (DAU)

-   The DAU is shown in FIG. 101. -   The DAU is composed of the following sub-blocks -   a. CPU Configuration and Arbitration Logic sub-block. -   b. Command Multiplexor sub-block. -   c. Read and Write Data Multiplexor sub-block.

The function of the DAU is to supply DRAM commands to the DCU.

-   -   The DCU requests a command from the DAU by asserting         dcu_dau_adv.     -   The DAU Command Multiplexor requests the Arbitration Logic         sub-block to arbitrate the next DRAM access. The Command         Multiplexor passes dcu_dau_adv as the re_arbitrate signal to the         Arbitration Logic sub-block.     -   If the RotationSync bit has been cleared, then the arbitration         logic grants exclusive access to the CPU and scheduled         refreshes. If the bit has been set, regular arbitration occurs.         A detailed description of RotationSync is given in section         20.14.12.2.1 on page 295.     -   Until the Arbitration Logic has a valid result it stalls the DCU         by asserting dau_dcu_msn2stall. The Arbitration Logic then         returns the selected arbitration winner to the Command         Multiplexor which issues the command to the DRAM. The         Arbitration Logic could stall for example if it selected a         shared read bus access but the Read Multiplexor indicated it was         busy by de-asserting read_cmd_rdy[1].     -   In the case of a read command the read data from the DRAM is         multiplexed back to the read requestor by the Read Multiplexor.         In the case of a write operation the Write Multiplexor         multiplexes the write data from the selected DIU write requester         to the DCU before the write command can occur. If the write data         is not available then the Command Multiplexor will keep         dau_dcu_valid de-asserted. This will stall the DCU until the         write command is ready to be issued.     -   Arbitration for non-CPU writes occurs in advance. The DCU         provides a signal dcu_dau_wadv which the Command Multiplexor         issues to the Arbitrate Logic as re_arbitrate_wadv. If         arbitration is blocked by the Write Multiplexor being busy, as         indicated by write_cmd_rdy[1] being de-asserted, then the         Arbitration Logic will stall the DCU by asserting         dau_dcu_msn2stall until the Write Multiplexor is ready.         20.14.10.1 Read Accesses

The timing of a non-CPU DIU read access are shown in FIG. 109. Note re_arbitrate is asserted in the MSN2 state of the previous access.

Note the fixed timing relationship between the read acknowledgment and the first rvalid for all non-CPU reads. This means that the second and any later reads in a back-to-back non-CPU sequence have their acknowledgments asserted one cycle later, i.e. in the “MSN1” DCU state. The timing of a CPU DIU read access is shown in FIG. 110. Note re_arbitrate is asserted in the MSN2 state of the previous access.

Some points can be noted from FIG. 109 and FIG. 110.

DIU requests:

-   -   For non-CPU accesses the <unit>_diu_rreq signals are registered         before the arbitration can occur.     -   For CPU accesses the cpu_diu_rreq signal is not registered to         reduce CPU DIU access latency.

Arbitration occurs when the dcu_dau_adv signal from the DCU is asserted. The DRAM address for the arbitration winner is available in the next cycle, the RST state of the DCU.

The DRAM access starts in the MSN1 state of the DCU and completes in the RST state of the DCU.

Read data is available:

-   -   In the MSN2 cycle where it is output unregistered to the CPU     -   In the MSN2 cycle and registered in the DAU before being output         in the next cycle to all other read requesters in order to ease         timing.

The DIU protocol is in fact:

-   -   Pipelined i.e. the following transaction is initiated while the         previous transfer is in progress.     -   Split transaction i.e. the transaction is split into independent         address and data transfers.

Some general points should be noted in the case of CPU accesses:

-   -   Since the CPU request is not registered in the DIU before         arbitration, then the CPU must generate the request, route it to         the DAU and complete arbitration all in 1 cycle. To facilitate         this CPU access is arbitrated late in the arbitration cycle (see         Section 20.14.12.2).     -   Since the CPU read data is not registered in the DAU and CPU         read data is available 8 ns after the start of the access then         4.5 ns are available for routing and any shallow logic before         the CPU read data is captured by the CPU (see Section 20.14.4).

The phases of CPU DIU read access are shown in FIG. 111. This matches the timing shown in Table 135.

20.14.10.2 Write Accesses

CPU writes are posted into a 1-deep write buffer in the DIU and written to DRAM as shown below in FIG. 112.

The sequence of events is as follows:

-   -   [1] The DIU signals that its buffer for CPU posted writes is         empty (and has been for some time in the case shown).     -   [2] The CPU asserts “cpu_diu_wdatavalid” to enable a write to         the DIU buffer and presents valid address, data and write mask.         The CPU considers the write posted and thus complete in the         cycle following [2] in the diagram below.     -   [3] The DIU stores the address/data/mask in its buffer and         indicates to the arbitration logic that a posted write wishes to         participate in any upcoming arbitration.     -   [4] Provided the CPU still has a pre-access entitlement left, or         is next in line for a round-robin award, a slot is arbitrated in         favour of the posted write. Note that posted CPU writes have         higher arbitration priority than simultaneous CPU reads.     -   [5] The DRAM write occurs.     -   [6] The earliest that “diu_cpu_write_rdy” can be re-asserted in         the “MSN1” state of the DRAM write. In the same cycle, having         seen the re-assertion, the CPU can asynchronously turn around         “cpu_diu_wdatavalid” and enable a subsequent posted write,         should it wish to do so.

The timing of a non-CPU/non-CDU DIU write access is shown below in FIG. 113.

Compared to a read access, write data is only available from the requester 4 cycles after the address. An extra cycle is used to ensure that data is first registered in the DAU, before being despatched to DRAM. As a result, writes are pre-arbitrated 5 cycles in advance of the main arbitration decision to actually write the data to memory.

The diagram above shows the following sequence of events:

-   -   [1] A non-CPU block signals a write request.     -   [2] A registered version of this is available to the DAU         arbitration logic.     -   [3] Write pre-arbitration occurs in favour of the requester.     -   [4] A write acknowledgment is returned by the DIU.     -   [5] The pre-arbitration will only be upheld if the requester         supplies 4 consecutive write data quarter-words, qualified by an         asserted wvalid flag.     -   [6] Provided this has happened, the main arbitration logic is in         a position at [6] to reconfirm the pre-arbitration decision.         Note however that such reconfirmation may have to wait a further         one or two DRAM accesses, if the write is pre-empted by a CPU         pre-access and/or a scheduled refresh.     -   [7] This is the earliest that the write to DRAM can occur.     -   Note that neither the arbitration at [8] nor the pre-arbitration         at [9] can award its respective slot to a non-CPU write, due to         the ban on back-to-back accesses.

The timing of a CDU DIU write access is shown overleaf in FIG. 114.

This is simular to a regular non-CPU write access, but uses page mode to carry out 4 consecutive DRAM writes to contiguous addresses. As a consequence, subsequent accesses are delayed by 6 cycles, as shown in the diagram. Note that a new write can be pre-arbitrated at [10] in FIG. 114.

20.14.11 Command Multiplexor Sub-Block

TABLE 136 Command Multiplexor Sub-block IO Definition Port name Pins I/O Description Clocks and Resets pclk 1 In System Clock prst_n 1 In System reset, synchronous active low DIU Read Interface to SoPEC Units <unit>_diu_(—) 17 In Read address to DIU radr[21:5] 17 bits wide (256-bit aligned word). diu_<unit>_(—) 1 Out Acknowledge from DIU that read rack request has been accepted and new read address can be placed on <unit>_diu_radr DIU Write Interface to SoPEC Units <unit>_diu_(—) 17 In Write address to DIU except CPU, wadr[21:5] SCB, CDU 17 bits wide (256-bit aligned word) cpu_diu_(—) 22 In CPU Write address to DIU wadr[21:4]] (128-bit aligned address.) cpu_diu_(—) 16 In Byte enables for CPU write. wmask cdu_diu_(—) 19 In CDU Write address to DIU wadr[21:3] 19 bits wide (64-bit aligned word) Addresses cannot cross a 256-bit word DRAM boundary. diu_<unit>_(—) 1 Out Acknowledge from DIU that write wack request has been accepted and new write address can be placed on <unit>_diu_wadr Outputs to CPU Interface and Arbitration Logic sub-block re_arbitrate 1 Out Signalling telling the arbitration logic to choose the next arbi- tration winner. re_arbitrate_(—) 1 Out Signal telling the arbitration wadv logic to choose the next arbi- tration winner for non-CPU writes 2 timeslots in advance Debug Outputs to CPU Configuration and Arbitration Logic Sub-block write_sel 5 Out Signal indicating the SoPEC Unit for which the current write trans- action is occurring. Encoding is described in Table . write_(—) 1 Out Signal indicating that write trans- complete action to SoPEC Unit indicated by write_sel is complete. Inputs from CPU Interface and Arbitration Logic sub-block arb_gnt 1 In Signal lasting 1 cycle which indi- cates arbitration has occurred and arb_sel is valid. arb_sel 5 In Signal indicating which requesting SoPEC Unit has won arbitration. Encoding is described in Table . dir_sel 2 In Signal indicating which sense of access associated with arb_sel 00: issue non-CPU write 01: read winner 10: write winner 11: refresh winner Inputs from Read Write Multiplexor Sub-block write_data_(—) 2 In Signal indicating that valid write valid data is available for the current command. 00 = not valid 01 = CPU write data valid 10 = non-CPU write data valid 11 = both CPU and non-CPU write data valid wdata 256 In 256-bit non-CPU write data cpu_wdata 32 In 32-bit CPU write data Outputs to Read Write Multiplexor Sub-block write_data_(—) 2 Out Signal indicating the Command accept Multiplexor has accepted the write data from the write multiplexor 00 = not valid 01 = accepts CPU write data 10 = accepts non-CPU write data 11 = not valid Inputs from DCU dcu_dau_(—) 1 In Signal indicating to DAU to adv supply next command to DCU dcu_dau_(—) 1 In Signal indicating to DAU to wadv initiate next non-CPU write Outputs to DCU dau_dcu_(—) 17 Out Signal indicating the address adr[21:5] for the DRAM access. This is a 256-bit aligned DRAM address. dau_dcu_(—) 1 Out Signal indicating the direction rwn for the DRAM access (1 = read, 0 = write). dau_dcu_(—) 1 Out Signal indicating if access is cduwpage a CDU write page mode access (1 = CDU page mode, 0 = not CDU page mode). dau_dcu_(—) 1 Out Signal indicating that a refresh refresh command is to be issued. If asserted dau_dcu_adr, dau_dcu_rwn and dau_dcu_cduwpage are ignored. dau_dcu_(—) 256 Out 256-bit write data to DCU wdata dau_dcu_(—) 32 Out Byte encoded write data mask for wmask 256-bit dau_dcu_wdata to DCU 20.14.11.1 Command Multiplexor Sub-Block Description

The Command Multiplexor sub-block issues read, write or refresh commands to the DCU, according to the SoPEC Unit selected for DRAM access by the Arbitration Logic. The Command Multiplexor signals the Arbitration Logic to perform arbitration to select the next SoPEC Unit for DRAM access. It does this by asserting the re_arbitrate signal. re_arbitrate is asserted when the DCU indicates on dcu_dau_adv that it needs the next command.

The Command Multiplexor is shown in FIG. 115.

Initially, the issuing of commands is described. Then the additional complexity of handling non-CPU write commands arbitrated in advance is introduced.

DAU-DCU Interface

See Section 20.14.5 for a description of the DAU-DCU interface.

Generating re_arbitrate

The condition for asserting re_arbitrate is that the DCU is looking for another command from the DAU. This is indicated by dcu_dau_adv being asserted. re_arbitrate=dcu_dau_adv Interface to SoPEC DIU Requestors

When the Command Multiplexor initiates arbitration by asserting re_arbitrate to the Arbitration Logic sub-block, the arbitration winner is indicated by the arb_sel[4:0] and dir_sel[1:0] signals returned from the Arbitration Logic. The validity of these signals is indicated by arb_gnt. The encoding of arb_sel[4:0] is shown in Table.

The value of arb_sel[4:0] is used to control the steering multiplexor to select the DIU address of the winning arbitration requestor. The arb_gnt signal is decoded as an acknowledge, diu_<unit>_*ack back to the winning DIU requestor. The timing of these operations is shown in FIG. 116. adr[21:0] is the output of the steering multiplexor controlled by arb_sel[4:0]. The steering multiplexor can acknowledge DIU requestors in successive cycles.

Command Issuing Logic

The address presented by the winning SoPEC requestor from the steering multiplexor is presented to the command issuing logic together with arb_sel[4:0] and dir_sel[1:0].

The command issuing logic translates the winning command into the signals required by the DCU. adr_(—)[21:0], arb_sel[4:0] and dir_sel[1:0] comes from the steering multiplexor.

dau_dcu_adr[21:5] = adr[21:5] dau_dcu_rwn = (dir_sel[1:0] == read) dau_dcu_cduwpage = (arb_sel[4:0] == CDU write) dau_dcu_refresh = (dir_sel[1:0]== refresh) dau_dcu_valid indicates that a valid command is available to the DCU.

For a write command, dau_dcu_valid will not be asserted until there is also valid write data present. This is indicated by the signal write_data_valid[1:0] from the Read Write Data Multiplexor sub-block.

For a write command, the data issued to the DCU on dau_dcu_wdata[255:0] is multiplexed from cpu_wdata[31:0] and wdata[255:0] depending on whether the write is a CPU or non-CPU write. The write data from the Write Multiplexor for the CDU is available on wdata[63:0]. This data must be issued to the DCU on dau_dcu_wdata[255:0]. wdata[63:0] is copied to each 64-bit word of dau_dcu_wdata[255:0].

dau_dcu_wdata[255:0] = 0x00000000 if (arb_sel[4:0]==CPU write) then dau_dcu_wdata[31:0] = cpu_wdata[31:0] elsif (arb_sel[4:0]==CDU write)) then dau_dcu_wdata[63:0] = wdata[63:0] dau_dcu_wdata[127:64] = wdata[63:0] dau_dcu_wdata[191:128] = wdata[63:0] dau_dcu_wdata[255:192] = wdata[63:0] else dau_dcu_wdata[255:0] = wdata[255:0] CPU Write Masking

The CPU write data bus is only 128 bits wide. cpu_diu_wmask[15:0] indicates how many bytes of that 128 bits should be written. The associated address cpu_diu_wadr[21:4] is a 128-bit aligned address. The actual DRAM write must be a 256-bit access. The command multiplexor issues the 256-bit DRAM address to the DCU on dau_dcu_adr[21:5]. cpu_diu_wadr[4] and cpu_diu_wmask[15:0] are used jointly to construct a byte write mask dau_dcu_wmask[31:0] for this 256-bit write access.

CDU Write Masking

The CPU performs four 64-bit word writes to 4 contiguous 256-bit DRAM addresses with the first address specified by cdu_diu_wadr[21:3]. The write address cdu_diu_wadr[21:5] is 256-bit aligned with bits cdu_diu_wadr[44:3] allowing the 64-bit word to be selected. If these 4 DRAM words lie in the same DRAM row then an efficient access will be obtained.

The command multiplexor logic must issue 4 successive accesses to 256-bit DRAM addresses cdu_diu_wadr[21:5],+1,+2,+3.

dau_dcu_wmask[31:0] indicates which 8 bytes (64-bits) of the 256-bit word are to be written.

dau_dcu_wmask[31:0] is calculated using cdu_diu_wadr[4:3] i.e. bits 8*cdu_diu_wadr[4:3] to 8*(cdu_diu_wadr[4:3]+1)−1 of dau_dcu_wmask[31:0]are asserted.

Arbitrating Non-CPU Writes in Advance

In the case of a non-CPU write commands, the write data must be transferred from the SoPEC requester before the write can occur. Arbitration should occur early to allow for any delay for the write data to be transferred to the DRAM.

FIG. 113 indicates that write data transfer over 64-bit busses will take a further 4 cycles after the address is transferred. The arbitration must therefore occur 4 cycles in advance of arbitration for read accesses, FIG. 109 and FIG. 110, or for CPU writes FIG. 112. Arbitration of CDU write accesses, FIG. 114, should take place 1 cycle in advance of arbitration for read and CPU write accesses. To simplify implementation CDU write accesses are arbitrated 4 cycles in advance, similar to other non-CPU writes.

The Command Multiplexor generates another version of re_arbitrate called re_arbitrate_wadv based on the signal dcu_dau_wadv from the DCU. In the 3 cycle DRAM access dcu_dau_adv and therefore re_arbitrate are asserted in the MSN2 state of the DCU state-machine. dcu_dau_wadv and therefore re_arbitrate_wadv will therefore be asserted in the following RST state, see FIG. 117. This matches the timing required for non-CPU writes shown in FIG. 113 and FIG. 114.

-   -   re_arbitrate_wadv causes the Arbitration Logic to perform an         arbitration for non-CPU in advance.

re_arbitrate = dcu_dau_adv re_arbitrate_wadv = dcu_dau_wadv

If the winner of this arbitration is a non-CPU write then arb_gnt is asserted and the arbitration winner is output on arb_sel[4:0] and dir_sel[1:0]. Otherwise arb_gnt is not asserted.

Since non-CPU write commands are arbitrated early, the non-CPU command is not issued to the DCU immediately but instead written into an advance command register.

if (arb_(—sel(4:0 == non-CPU write) then) advance_cmd_register[3:0] = arb_sel[4:0] advance_cmd_register[5:4] = dir_sel[1:0] advance_cmd_register[27:6] = adr[21:0]

If a DCU command is in progress then the arbitration in advance of a non-CPU write command will overwrite the steering multiplexor input to the command issuing logic. The arbitration in advance happens in the DCU MSN1 state. The new command is available at the steering multiplexor in the MSN2 state. The command in progress will have been latched in the DRAM by MSN falling at the start of the MSN1 state.

Issuing Non-CPU Write Commands

The arb_sel[4:0] and dir_(—)[1:0] values generated by the Arbitration Logic reflect the out of order arbitration sequence.

This out of order arbitration sequence is exported to the Read Write Data Multiplexor sub-block. This is so that write data in available in time for the actual write operation to DRAM. Otherwise a latency would be introduced every time a write command is selected.

However, the Command Multiplexor must execute the command stream in-order. In-order command execution is achieved by waiting until re_arbitrate has advanced to the non-CPU write timeslot from which re_arbitrate_wadv has previously issued a non-CPU write written to the advance command register.

If re_arbitrate_wadv arbitrates a non-CPU write in advance then within the Arbitration Logic the timeslot is marked to indicate whether a write was issued.

When re_arbitrate advances to a write timeslot in the Arbitration Logic then one of two actions can occur depending on whether the slot was marked by re_arbitrate_wadv to indicate whether a write was issued or not.

-   -   Non-CPU write arbitrated by re_arbitrate_wadv

If the timeslot has been marked as having issued a write then the arbitration logic responds to re_arbitrate by issuing arb_sel[4:0], dir_sel[1:0] and asserting arb_gnt as for a normal arbitration but selecting a non-CPU write access. Normally, re_arbitrate does not issue non-CPU write accesses. Non-CPU writes are arbitrated by re_arbitrate_wadv. dir_sel[1:0]==00 indicates a non-CPU write issued by re_arbitrate.

The command multiplexor does not write the command into the advance command register as it has already been placed there earlier by re_arbitrate_wadv. Instead, the already present write command in the advance command register is issued when write_data_valid[1]=1. Note, that the value of arb_sel[4:0] issued by re_arbitrate could specify a different write than that in the advance command register since time has advanced. It is always the command in the advance command register that is issued. The steering multiplexor in this case must not issue an acknowledge back to SoPEC requester indicated by the value of arb_sel[4:0].

if (dir_sel[1:0] == 00) then command_issuing_logic[27:0] == advance_cmd_register[27:0] else command_issuing_logic[27:0] == steering_multiplexor[27:0] ack = arb_gnt AND NOT (dir_sel[1:0] == 00)

-   -   Non-CPU write not arbitrated by re_arbitrate_wadv

If the timeslot has been marked as not having issued a write, the re_arbitrate will use the un-used read timeslot selection to replace the un-used write timeslot with a read timeslot according to Section 20.10.6.2 Unused read timeslots allocation.

The mechanism for write timeslot arbitration selects non-CPU writes in advance. But the selected non-CPU write is stored in the Command Multiplexor and issued when the write data is available. This means that even if this timeslot is overwritten by the CPU reprogramming the timeslot before the write command is actually issued to the DRAM, the originally arbitrated non-CPU write will always be correctly issued.

Accepting Write Commands

When a write command is issued then write_data_accept[1:0] is asserted. This tells the Write Multiplexor that the current write data has been accepted by the DRAM and the write multiplexor can receive write data from the next arbitration winner if it is a write. write_data_accept[1:0] differentiates between CPU and non-CPU writes. A write command is known to have been issued when re_arbitrate_wadv to decide on the next command is detected.

In the case of CDU writes the DCU will generate a signal dcu_dau_cduwaccept which tells the Command Multiplexor to issue a write_data_accept[1]. This will result in the Write Multiplexor supplying the next CDU write data to the DRAM.

write_data_accept[0] = RISING EDGE(re_arbitrate_wadv) AND command_issuing_logic(dir_sel[1]==1) AND command_issuing_logic(arb_sel[4:0]==CPU) write_data_accept[1] = (RISING EDGE(re_arbitrate_wadv) AND command_issuing_logic(dir_sel[1]==1) AND command_issuing_logic(arb_sel[4:0]==non_CPU)) OR dcu_dau_cduwaccept==1

Debug logic output to CPU Configuration and Arbitration Logic sub-block write_sel[4:0] reflects the value of arb_sel[4:0] at the command issuing logic. The signal write complete is asserted when every any bit of write_data_accept[1:0] is asserted.

write_complete = write_data_accept[0] OR write_data_accept[0] write_sel[4:0] and write_complete are CPU readable from the DIUPerformance nad WritePerformance status registers. When write_complete is asserted write_sel[4:0] will indicate which write access the DAU has issued. 20.14.12 CPU Configuration and Arbitration Logic Sub-Block

TABLE 137 CPU Configuration and Arbitration Logic Sub-block IO Definition Port name Pins I/O Description Clocks and Resets Pclk 1 In System Clock prst_n 1 In System reset, synchronous active low CPU Interface data and control signals cpu_adr[10:2] 9 In 9 bits (bits 10:2) are required to decode the configuration register address space. cpu_dataout 32 In Shared write data bus from the CPU for DRAM and configuration data diu_cpu_data 32 Out Configuration, status and debug read data bus to the CPU diu_cpu_(—) 1 Out Signal indicating the data on debug_valid the diu_cpu_data bus is valid debug data. cpu_rwn 1 In Common read/not-write signal from the CPU cpu_acode 2 In CPU access code signals. cpu_acode[0] - Program (0)/ Data (1) access cpu_acode[1] - User (0)/ Supervisor (1) access The DAU will only allow supervisor mode accesses to data space. cpu_diu_sel 1 In Block select from the CPU. When cpu_diu_sel is high both cpu_adr and cpu_dataout are valid diu_cpu_rdy 1 Out Ready signal to the CPU. When diu_cpu_rdy is high it indicates the last cycle of the access. For a write cycle this means cpu_dataout has been registered by the block and for a read cycle this means the data on diu_cpu_data is valid. diu_cpu_berr 1 Out Bus error signal to the CPU indicating an invalid access. DIU Read Interface to SoPEC Units <unit>_diu_(—) 11 In SoPEC unit requests DRAM read. rreq DIU Write Interface to SoPEC Units diu_cpu_(—) 1 In Indicator that CPU posted write write_rdy buffer is empty. <unit>_diu_(—) 4 In Non- CPU SoPEC unit requests wreq DRAM write. Inputs from Command Multiplexor sub-block re_arbitrate 1 In Signal telling the arbitration logic to choose the next arbitration winner. re_arbitrate_(—) 1 In Signal telling the arbitration wadv logic to choose the next arbitration winner for non-CPU writes 2 timeslots in advance Outputs to DCU dau_dcu_(—) 1 Out Signal indicating from DAU msn2stall Arbitration Logic which when asserted stalls DCU in MSN2 state. Inputs from Read and Write Multiplexor sub-block read_cmd_rdy 2 In Signal indicating that read multiplexor is ready for next read read command. 00 = not ready 01 = ready for CPU read 10 = ready for non-CPU read 11 = ready for both CPU and non-CPU reads write_cmd_rdy 2 In Signal indicating that write multiplexor is ready for next write command. 00 = not ready 01 = ready for CPU write 10 = ready for non-CPU write 11 = ready for both CPU and non-CPU write Outputs to other DAU sub-blocks arb_gnt 1 In Signal lasting 1 cycle which indicates arbitration has occurred and arb_sel is valid. arb_sel 5 In Signal indicating which requesting SoPEC Unit has won arbitration. Encoding is described in Table . dir_sel 2 In Signal indicating which sense of access associated with arb_sel 00: issue non-CPU write 01: read winner 10: write winner 11: refresh winner Debug Inputs from Read-Write Multiplexor sub-block read_sel 5 In Signal indicating the SoPEC Unit for which the current read trans- action is occurring. Encoding is described in Table . read_complete 1 In Signal indicating that read transaction to SoPEC Unit indicated by read_sel is complete. Debug Inputs from Command Multiplexor sub-block write_sel 5 In Signal indicating the SoPEC Unit for which the current write transaction is occurring. Encoding is described in Table . write_complete 1 In Signal indicating that write transaction to SoPEC Unit indicated by write_sel is complete. Debug Inputs from DCU dcu_dau_(—) 1 In Signal indicating that the DCU refreshcomplete has completed a refresh. Debug Inputs from DAU IO various n In Various DAU IO signals which can be monitored in debug mode

The CPU Interface and Arbitration Logic sub-block is shown in FIG. 118.

20.14.12.1 CPU Interface and Configuration Registers Description

The CPU Interface and Configuration Registers sub-block provides for the CPU to access DAU specific registers by reading or writing to the DAU address space.

The CPU subsystem bus interface is described in more detail in Section 11.4.3. The DAU block will only allow supervisor mode accesses to data space (i.e. cpu_acode[1:0]=b11). All other accesses will result in diu_cpu_berr being asserted.

The configuration registers described in Section 20.14.9

TABLE 130 DAU configuration registers Address # (DIU_base+) Register bits Reset Description Reset 0x00 Reset 1 0x1 A write to this register causes a reset of the DIU. This register can be read to indicate the reset state: 0 - reset in progress 1 - reset not in progress Refresh 0x04 RefreshPeriod 9 0x063 Refresh controller. When set to 0 refresh is off, otherwise the value indicates the number of cycles, less one, between each refresh. [Note that for a system clock fre- quency of 160 MHz, a value exceeding 0x63 (indicating a 100-cycle refresh period) should not be programmed, or the DRAM will malfunction.] Timeslot allocation and control 0x08 NumMainTimeslots 6 0x01 Number of main timeslots (1-64) less one 0x0C CPUPreAccessTimes 4 0x0 (CPUPreAccessTimeslots + 1) lots main slots out of a total of (CPUTotalTimeslots + 1) are preceded by a CPU access. 0x10 CPUTotalTimeslots 4 0x0 (CPUPreAccessTimeslots + 1) main slots out of a total of (CPUTotalTimeslots + 1) are preceded by a CPU access. 0x100-0x1FC MainTimeslot[63:0] 64 × 4 [63:1][3:0] = 0x0 Programmable main timeslots [0][3:0] = 0xE (up to 64 main timeslots). 0x200 ReadRoundRobinLevel 12 0x000 For each read requester plus refresh 0 = level1 of round-robin 1 = level2 of round-robin The bit order is defined in Table . 0x204 EnableCPURoundRobin 1 0x1 Allows the CPU to particpate in the unused read round-robin scheme. If disabled, the shared CPU/refresh round-robin position is dedicated solely to refresh. 0x208 RotationSync 1 0x1 Writing 0, followed by 1 to this bit allows the timeslot rotation to advance on a cycle basis which can be determined by the CPU. 0x20C minNonCPUReadAdr 12 0x800 12 MSBs of lowest DRAM address which may be read by non-CPU requesters. 0x210 minDWUWriteAdr 12 0x800 12 MSBs of lowest DRAM address which may be written to by the DWU. 0x214 minNonCPUWriteAdr 12 0x800 12 MSBs of lowest DRAM address which may be written to by non-CPU requesters other than the DWU. Debug 0x300 DebugSelect[11:2] 10 0x304 Debug address select. Indicates the address of the register to report on the diu_cpu_data bus when it is not otherwise being used. When this signal carries debug information the signal diu_cpu_debug_valid will be asserted. Debug: arbitration and performance 0x304 ArbitrationHistory 22 — Bit 0 = arb_gnt Bit 1 = arb_executed Bit 6:2 = arb_sel[4:0] Bit 12:7 = timeslot_number[5:0] Bit 15:13 = access_type[2:0] Bit 16 = back2back_non_cpu_write Bit 17 = sticky_back2back_non_cpu_write (Sticky version of same, cleared on reset.) Bit 18 = rotation_sync Bit 20:19 = rotation_state Bit 21 = sticky_invalid_non_cpu_adr See Section 20.14.9.2 DIU Debug for a description of the fields. Read only register. 0x308 DIUPerformance 31 — Bit 0 = cpu_diu_rreq Bit 1 = scb_diu_rreq Bit 2 = cdu_diu_rreq Bit 3 = cfu_diu_rreq Bit 4 = lbd_diu_rreq Bit 5 = sfu_diu_rreq Bit 6 = td_diu_rreq Bit 7 = tfs_diu_rreq Bit 8 = hcu_diu_rreq Bit 9 = dnc_diu_rreq Bit 10 = llu_diu_rreq Bit 11 = pcu_diu_rreq Bit 12 = cpu_diu_wreq Bit 13 = scb_diu_wreq Bit 14 = cdu_diu_wreq Bit 15 = sfu_diu_wreq Bit 16 = dwu_diu_wreq Bit 17 = refresh_req Bit 22:18 = read_sel[4:0] Bit 23 = read_complete Bit 28:24 = write_sel[4:0] Bit 29 = write_complete Bit 30 = dcu_dau_refreshcomplete See Section 20.14.9.2 DIU Debug for a description of the fields. Read only register. Debug DIU read requesters interface signals 0x30C CPUReadInterface 25 — Bit 0 = cpu_diu_rreq Bit 22:1 = cpu_adr[21:0] Bit 23 = diu_cpu_rack Bit 24 = diu_cpu_rvalid Read only register. 0x310 SCBReadInterface 20 Bit 0 = scb_diu_rreq Bit 17:1 = scb_diu_radr[21:5] Bit 18 = diu_scb_rack Bit 19 = diu_scb_rvalid Read only register. 0x314 CDUReadInterface 20 — Bit 0 = cdu_diu_rreq Bit 17:1 = cdu_diu_radr[21:5] Bit 18 = diu_cdu_rack Bit 19 = diu_cdu_rvalid Read only register. 0x318 CFUReadInterface 20 — Bit 0 = cfu_diu_rreq Bit 17:1 = cfu_diu_radr[21:5] Bit 18 = diu_cfu_rack Bit 19 = diu_cfu_rvalid Read only register. 0x31C LBDReadInterface 20 — Bit 0 = lbd_diu_rreq Bit 17:1 = lbd_diu_radr[21:5] Bit 18 = diu_lbd_rack Bit 19 = diu_lbd_rvalid Read only register. 0x320 SFUReadInterface 20 — Bit 0 = sfu_diu_rreq Bit 17:1 = sfu_diu_radr[21:5] Bit 18 = diu_sfu_rack Bit 19 = diu_sfu_rvalid Read only register. 0x324 TDReadInterface 20 — Bit 0 = td_diu_rreq Bit 17:1 = td_diu_radr[21:5] Bit 18 = diu_td_rack Bit 19 = diu_td_rvalid Read only register. 0x328 TFSReadInterface 20 — Bit 0 = tfs_diu_rreq Bit 17:1 = tfs_diu_radr[21:5] Bit 18 = diu_tfs_rack Bit 19 = diu_tfs_rvalid Read only register. 0x32C HCUReadInterface 20 — Bit 0 = hcu_diu_rreq Bit 17:1 = hcu_diu_radr[21:5] Bit 18 = diu_hcu_rack Bit 19 = diu_hcu_rvalid Read only register. 0x330 DNCReadInterface 20 — Bit 0 = dnc_diu_rreq Bit 17:1 = dnc_diu_radr[21:5] Bit 18 = diu_dnc_rack Bit 19 = diu_dnc_rvalid Read only register. 0x334 LLUReadInterface 20 — Bit 0 = llu_diu_rreq Bit 17:1 = lluu_diu_radr[21:5] Bit 18 = diu_llu_rack Bit 19 = diu_llu_rvalid Read only register. 0x338 PCUReadInterface 20 — Bit 0 = pcu_diu_rreq Bit 17:1 = pcu_diu_radr[21:5] Bit 18 = diu_pcu_rack Bit 19 = diu_pcu_rvalid Read only register. Debug DIU write requesters interface signals 0x33C CPUWriteInterface 27 — Bit 0 = cpu_diu_wreq Bit 22:1 = cpu_adr[2.1:0] Bit 24:23 = cpu_diu_wmask[1:0] Bit 25 = diu_cpu_wack Bit 26 = cpu_diu_wvalid Read only register. 0x340 SCBWriteInterface 20 — Bit 0 = scb_diu_wreq Bit 17:1 = scb_diu_wadr[21:5] Bit 18 = diu_scb_wack Bit 19 = scb_diu_wvalid Read only register. 0x344 CDUWriteInterface 22 — Bit 0 = cdu_diu_wreq Bit 19:1 = cdu_diu_wadr[21:3] Bit 20 = diu_cdu_wack Bit 21 = cdu_diu_wvalid Read only register. 0x348 SFUWriteInterface 20 — Bit 0 = sfu_diu_wreq Bit 17:1 = sfu_diu_wadr[21:5] Bit 18 = diu_sfu_wack Bit 19 = sfu_diu_wvalid Read only register. 0x34C DWUWriteInterface 20 — Bit 0 = dwu_diu_wreq Bit 17:1 = dwu_diu_wadr[21:5] Bit 18 = diu_dwu_wack Bit 19 = dwu_diu_wvalid Read only register. Debug DAU-DCU interface signals 0x350 DAU-DCUInterface 25 — Bit 16:0 = dau_dcu_adr[21:5] Bit 17 = dau_dcu_rwn Bit 18 = dau_dcu_cduwpage Bit 19 = dau_dcu_refresh Bit 20 = dau_dcu_msn2stall Bit 21 = dcu_dau_adv Bit 22 = dcu_dau_wadv Bit 23 = dcu_dau_refreshcomplete Bit 24 = dcu_dau_rvalid Read only register. are implemented here. 20.14.12.2 Arbitration Logic Description

Arbitration is triggered by the signal re_arbitrate from the Command Multiplexor sub-block with the signal arb_gnt indicating that arbitration has occurred and the arbitration winner is indicated by arb_sel[4:0]. The encoding of arb_sel[4:0] is shown in Table. The signal dir_sel[1:0] indicates if the arbitration winner is a read, write or refresh. Arbitration should complete within one clock cycle so arb_gnt is normally asserted the clock cycle after re_arbitrate and stays high for 1 clock cycle. arb_sel[4:0] and dir_sel[1:0] remain persistent until arbitration occurs again. The arbitration timing is shown in FIG. 119.

20.14.12.2.1 Rotation Synchronisation

A configuration bit, RotationSync, is used to initialise advancement through the timeslot rotation, in order that the CPU will know, on a cycle basis, which timeslot is being arbitrated. This is essential for debug purposes, so that exact arbitration sequences can be reproduced.

In general, if RotationSync is set, slots continue to be arbitrated in the regular order specified by the timeslot rotation. When the bit is cleared, the current rotation continues until the slot pointers for pre- and main arbitration reach zero. The arbitration logic then grants DRAM access exclusively to the CPU and refreshes.

When the CPU again writes to RotationSync to cause a 0-to-1 transition of the bit, the rdy acknowledgment back to the CPU for this write will be exactly coincident with the RST cycle of the initial refresh which heralds the enabling of a new rotation. This refresh, along with the second access which can be either a CPU pre-access or a refresh, (depending on the CPU's request inputs), form a 2-access “preamble” before the first non-CPU requester in the new rotation can be serviced. This preamble is necessary to give the write pre-arbitration the necessary head start on the main arbitration, so that write data can be loaded in time. See FIG. 105 below. The same preamble procedure is followed when emerging from reset.

The alignment of rdy with the commencement of the rotation ensures that the CPU is always able to calculate at any point how far a rotation has progressed. RotationSync has a reset value of 1 to ensure that the default power-up rotation can take place.

Note that any CPU writes to the DIU's other configuration registers should only be made when RotationSync is cleared. This ensures that accesses by non-CPU requesters to DRAM are not affected by partial configuration updates which have yet to be completed.

20.14.12.2.2 Motivation for Rotation Synchronisation

The motivation for this feature is that communications with SoPEC from external sources are synchronised to the internal clock of our position within a DIU full timeslot rotation. This means that if an external source told SoPEC to start a print 3 separate times, it would likely be at three different points within a full DIU rotation. This difference means that the DIU arbitration for each of the runs would be different, which would manifest itself externally as anomalous or inconsistent print performance. The lack of reproducibility is the problem here.

However, if in response to the external source saying to start the print, we caused the internal to pass through a known state at a fixed time offset to other internal actions, this would result in reproducible prints. So, the plan is that the software would do a rotation synchronise action, then writes “Go” into various PEP units to cause the prints. This means the DIU state will be the identical with respect to the PEP units state between separate runs.

20.14.12.2.3 Wind-Down Protocol When Rotation Synchronisation is Initiated

When a zero is written to “RotationSync”, this initiates a “wind-down protocol” in the DIU, in which any rotation already begun must be fully completed. The protocol implements the following sequence:—

-   -   The pre-arbitration logic must reach the end of whatever         rotation it is on and stop pre-arbitrating.     -   Only when this has happened, does the main arbitration consider         doing likewise with its current rotation. Note that the main         arbitration lags the pre-arbitration by at least 2 DRAM         accesses, subject to variation by CPU pre-accesses and/or         scheduled refreshes, so that the two arbitration processes are         sometimes on different rotations.     -   Once the main arbitration has reached the end of its rotation,         rotation synchronisation is considered to be fully activated.         Arbitration then proceeds as outlined in the next section.         20.14.12.2.4 Arbitration During Rotation Synchronisation

Note that when RotationSync is ‘0’ and, assuming the terminating rotation has completely drained out, then DRAM arbitration is granted according to the following fixed priority order:—Scheduled Refresh->CPU(W)->CPU(R)->Default Refresh.

CPU pre-access counters play no part in arbitration during this period. It is only subsequently, when emerging from rotation sync, that they are reloaded with the values of CPUPreAccessTimeslots and CPUTotalTimeslots and normal service resumes.

20.14.12.2.5 Timeslot-Based Arbitration

Timeslot-based arbitration works by having a pointer point to the current timeslot. This is shown in FIG. 95 repeated here as FIG. 121. When re-arbitration is signaled the arbitration winner is the current timeslot and the pointer advances to the next timeslot. Each timeslot denotes a single access. The duration of the timeslot depends on the access.

If the SoPEC Unit assigned to the current timeslot is not requesting then the unused timeslot arbitration mechanism outlined in Section 20.10.6 is used to select the arbitration winner. Note that this unused slot re-allocation is guaranteed to produce a result, because of the inclusion of refresh in the round-robin scheme.

Pseudo-code to represent arbitration is given below:

if re_arbitrate == 1 then arb_gnt = 1 if current timeslot requesting then choose (arb_sel, dir_sel) at current timeslot else // un-used timeslot scheme choose winner according to un-used timeslot allocation of Section 20.10.6 arb_gnt = 0 20.14.12.3 Arbitrating Non-CPU Writes in Advance

In the case of a non-CPU write commands, the write data must be transferred from the SoPEC requester before the write can occur. Arbitration should occur early to allow for any delay for the write data to be transferred to the DRAM.

FIG. 113 indicates that write data transfer over 64-bit busses will take a further 4 cycles after the address is transferred. The arbitration must therefore occur 4 cycles in advance of arbitration for read accesses, FIG. 109 and FIG. 110, or for CPU writes FIG. 112. Arbitration of CDU write accesses, FIG. 114, should take place 1 cycle in advance of arbitration for read and CPU write accesses. To simplify implementation CDU write accesses are arbitrated 4 cycles in advance, similar to other non-CPU writes.

The Command Multiplexor generates a second arbitration signal re_arbitrate_wadv which initiates the arbitration in advance of non-CPU write accesses.

The timeslot scheme is then modified to have 2 separate pointers:

-   -   re_arbitrate can arbitrate read, refresh and CPU read and write         accesses according to the position of the current timeslot         pointer.     -   re_arbitrate_wadv can arbitrate only non-CPU write accesses         according to the position of the write lookahead pointer.         -   Pseudo-code to represent arbitration is given below:

//re_arbitrate if (re_arbitrate == 1) AND (current timeslot pointer!= non- CPU write) then arb_gnt = 1 if current timeslot requesting then choose(arb_sel, dir_sel) at current timeslot else // un-used read timeslot scheme choose winner according to un-used read timeslot allocation of Section 20.10.6.2

If the SoPEC Unit assigned to the current timeslot is not requesting then the unused read timeslot arbitration mechanism outlined in Section 20.10.6.2 is used to select the arbitration winner.

//re_arbitrate_wadv if (re_arbitrate_wadv == 1) AND (write lookahead timeslot pointer == non-CPU write) then if write lookahead timeslot requesting then choose (arb_sel, dir_sel) at write lookahead timeslot arb_gnt = 1 elsif un-used write timeslot scheme has a requestor choose winner according to un-used write timeslot allocation of Section 20.10.6.1 arb_gnt = 1 else //no arbitration winner arb_gnt = 0

-   -   re_arbitrate is generated in the MSN2 state of the DCU         state-machine, whereas     -   re_arbitrate_wadv is generated in the RST state. See FIG. 103.

The write lookahead pointer points two timeslots in advance of the current timeslot pointer. Therefore re_arbitrate_wadv causes the Arbitration Logic to perform an arbitration for non-CPU two timeslots in advance. As noted in Table, each timeslot lasts at least 3 cycles. Therefor re_arbitrate_wadv arbitrates at least 4 cycles in advance.

At initialisation, the write lookahead pointer points to the first timeslot. The current timeslot pointer is invalid until the write lookahead pointer advances to the third timeslot when the current timeslot pointer will point to the first timeslot. Then both pointers advance in tandem.

Some accesses can be preceded by a CPU access as in Table. These CPU accesses are not allocated timeslots. If this is the case the timeslot will last 3 (CPU access)+3 (non-CPU access) =6 cycles. In that case, a second write lookahead pointer, the CPU pre-access write lookahead pointer, is selected which points only one timeslot in advance. re_arbitrate_wadv will still arbitrate 4 cycles in advance.

20.14.12.3.1 Issuing Non-CPU Write Commands

Although the Arbitration Logic will arbitrate non-CPU writes in advance, the Command Multiplexor must issue all accesses in the timeslot order. This is achieved as follows:

If re_arbitrate_wadv arbitrates a non-CPU write in advance then within the Arbitration Logic the timeslot is marked to indicate whether a write was issued.

//re_arbitrate_wadv if (re_arbitrate_wadv == 1) AND (write lookahead timeslot pointer == non-CPU write) then if write lookahead timeslot requesting then choose (arb_sel, dir_sel) at write lookahead timeslot arb_gnt = 1 MARK_timeslot = 1 elsif un-used write timeslot scheme has a requestor choose winner according to un-used write timeslot allocation of Section 20.10.6.1 arb_gnt = 1 MARK_timeslot = 1 else //no pre-arbitration winner arb_gnt = 0 MARK_timeslot = 0

When re_arbitrate advances to a write timeslot in the Arbitration Logic then one of two actions can occur depending on whether the slot was marked by re_arbitrate_wadv to indicate whether a write was issued or not.

-   -   Non-CPU write arbitrated by re_arbitrate_wadv

If the timeslot has been marked as having issued a write then the arbitration logic responds to re_arbitrate by issuing arb_sel[4:0], dir_sel[1:0] and asserting arb gnt as for a normal arbitration but selecting a non-CPU write access. Normally, re_arbitrate does not issue non-CPU write accesses. Non-CPU writes are arbitrated by re_arbitrate_wadv. dir_sel[1:0]==00 indicates a non-CPU write issued by re_arbitrate.

-   -   Non-CPU write not arbitrated by re_arbitrate_wadv

If the timeslot has been marked as not having issued a write, the re_arbitrate will use the un-used read timeslot selection to replace the un-used write timeslot with a read timeslot according to Section 20.10.6.2 Unused read timeslots allocation.

//re_arbitrate except for non-CPU writes if (re_arbitrate == 1) AND (current timeslot pointer!= non- CPU write) then arb_gnt = 1 if current timeslot requesting then choose(arb_sel, dir_sel) at current timeslot else // un-used read timeslot scheme choose winner according to un-used read timeslot allocation of Section 20.10.6.2 arb_gnt = 1 //non-CPU write MARKED as issued elsif (re_arbitrate == 1) AND (current timeslot pointer == non-CPU write) AND (MARK_timeslot == 1) then //indicate to Command Multiplexor that non-CPU write has been arbitrated in //advance arb_gnt = 1 dir_sel[1:0] == 00 //non-CPU write not MARKED as issued elsif (re_arbitrate == 1) AND (current timeslot pointer == non-CPU write) AND (MARK_timeslot == 0) then choose winner according to un-used read timeslot allocation of Section 20.10.6.2 arb_gnt = 1 20.14.12.4 Flow Control

If read commands are to win arbitration, the Read Multiplexor must be ready to accept the read data from the DRAM. This is indicated by the read_cmd_rdy[1:0] signal. read_cmd_rdy[1:0] supplies flow control from the Read Multiplexor.

read_cmd_rdy[0]==1 //Read multiplexor ready for CPU read read_cmd_rdy[1]==1 //Read multiplexor ready for non-CPU read

The Read Multiplexor will normally always accept CPU reads, see Section 20.14.13.1, so read_cmd_rdy[0]==1 should always apply.

Similarly, if write commands are to win arbitration, the Write Multiplexor must be ready to accept the write data from the winning SoPEC requestor. This is indicated by the write_cmd_rdy[1:0] signal. write_cmd_rdy[1:0] supplies flow control from the Write Multiplexor.

write_cmd_rdy[0]==1 //Write multiplexor ready for CPU write write_cmd_rdy[1]==1 //Write multiplexor ready for non- CPU write

The Write Multiplexor will normally always accept CPU writes, see Section 20.14.13.2, so write_cmd_rdy[0]==1 should always apply.

Non-CPU Read Flow Control

If re_arbitrate selects an access then the signal dau_dcu_msn2stall is asserted until the Read Write Multiplexor is ready.

arb_gnt is not asserted until the Read Write Multiplexor is ready.

This mechanism will stall the DCU access to the DRAM until the Read Write Multiplexor is ready to accept the next data from the DRAM in the case of a read.

//other access flow control dau_dcu_msn2stall = (((re_arbitrate selects CPU read) AND read_cmd rdy[0]==0) OR (re_arbitrate selects non-CPU read) AND read_cmd_rdy[1]==0)) arb_gnt not asserted until dau_dcu_msn2stall de-asserts 20.14.12.5 Arbitration Hierarchy

CPU and refresh are not included in the timeslot allocations defined in the DAU configuration registers of Table.

The hierarchy of arbitration under normal operation is

-   a. CPU access -   b. Refresh access -   c. Timeslot access.

This is shown in FIG. 124. The first DRAM access issued after reset must be a refresh.

As shown in FIG. 118, the DIU request signals <unit>_diu_rreq, <unit>_diu_wreq are registered at the input of the arbitration block to ease timing. The exceptions are the refresh_req signal, which is generated locally in the sub-block and cpu_diu_rreq. The CPU read request signal is not registered so as to keep CPU DIU read access latency to a minimum. Since CPU writes are posted, cpu_diu_wreq is registered so that the DAU can process the write at a later juncture. The arbitration logic is coded to perform arbitration of non-CPU requests first and then to gate the result with the CPU requests. In this way the CPU can make the requests available late in the arbitration cycle.

Note that when RotationSync is set to ‘0’, a modified hierarchy of arbitration is used. This is outlined in section 20.14.12.2.3 on page 280.

20.14.12.6 Timeslot Access

The basic timeslot arbitration is based on the MainTimeslot configuration registers. Arbitration works by the timeslot pointed to by either the current or write lookahead pointer winning arbitration. The pointers then advance to the next timeslot. This was shown in FIG. 90.

Each main timeslot pointer gets advanced each time it is accessed regardless of whether the slot is used.

20.14.12.7 Unused Timeslot Allocation

If an assigned slot is not used (because its corresponding SoPEC Unit is not requesting) then it is reassigned according to the scheme described in Section 20.10.6.

Only used non-CPU accesses are reallocated. CDU write accesses cannot be included in the unused timeslot allocation for write as CDU accesses take 6 cycles. The write accesses which the CDU write could otherwise replace require only 3 or 4 cycles.

Unused write accesses are re-allocated according to the fixed priority scheme of Table. Unused read timeslots are re-allocated according to the two-level round-robin scheme described in Section 20.10.6.2.

A pointer points to the most recently re-allocated unit in each of the round-robin levels. If the unit immediately succedling the pointer is requesting, then this unit wins the arbitration and the pointer is advanced to reflect the new winner. If this is not the case, then the subsequent units (wrapping back eventually to the pointed unit) in the level 1 round-robin are examined. When a requesting unit is found this unit wins the arbitration and the pointer is adjusted. If no unit is requesting then the pointer does not advance and the second level of round-robin is examined in a similar fashion. In the following pseudo-code the bit indices are for the ReadRoundRobinLevel configuration register described in Table.

//choose the winning arbitration level level1 = 0 level2 = 0 for i = 0 to 11 if unit(i) requesting AND ReadRoundRobinLevel(i) = 0 then level1 = 1 if unit(i) requesting AND ReadRoundRobinLevel(i) = 1 then level2 = 1

Round-robin arbitration is effectively a priority assignment with the units assigned a priority according to the round-robin order of Table but starting at the unit currently pointed to.

//levelptr is pointer of selected round robin level priority is array 0 to 11 // index 0 is SCBR(0) etc. from Table //assign decreasing priorities from the current pointer; maximum priority is 11 for i = 1 to 12 priority (levelptr + i) = 12 − i i++

The arbitration winner is the one with the highest priority provided it is requesting and its ReadRoundRobinLevel bit points to the chosen level. The levelptr is advanced to the arbitration winner.

The priority comparison can be done in the hierarchical manner shown in FIG. 125.

20.14.12.8 How Non-CPU Address Restrictions Affect Arbitration

Recall from Table “DAU configuration registers,” on page 288, “DAU configuration registers,” on page 268 that there are minimum valid DRAM addresses for non-CPU accesses, defined by minNonCPUReadAdr, minDWUWriteAdr and minNonCPUWriteAdr. Similarly, a non-CPU requester may not try to access a location above the high memory mark.

To ensure compliance with these address restrictions, the following DIU response occurs for any incorrectly addressed non-CPU writes:—

-   -   Issue a write acknowledgment at pre-arbitration time, to prevent         the write requester from hanging.     -   Disregard the incoming write data and write valids and void the         pre-arbitration.     -   Subsequently re-allocate the write slot at main arbitration time         via the round robin.

For any incorrectly addressed non-CPU reads, the response is:—

-   -   Arbitrate the slot in favour of the scheduled, misbehaving         requester.     -   Issue the read acknowledgement and rvalids to keep the requester         from hanging.     -   Intercept the read data coming from the DCU and send back all         zeros instead.

If an invalidly addressed non-CPU access is attempted, then a sticky bit, sticky_invalid_non_cpu_adr, is set in the ArbitrationHistory configuration register. See Table n page 293 on page 275 for details.

20.14.12.9 Refresh Controller Description

The refresh controller implements the functionality described in detail in Section 20.10.5. Refresh is not included in the timeslot allocations.

CPU and refresh have priority over other accesses. If the refresh controller is requesting i.e. refresh_req is asserted, then the refresh request will win any arbitration initiated by re_arbitrate. When the refresh has won the arbitration refresh_req is de-asserted.

The refresh counter is reset to RefreshPeriod[8:0] i.e. the number of cycles between each refresh. Every time this counter decrements to 0, a refresh is issued by asserting refresh_req. The counter immediately reloads with the value in RefreshPeriod[8:0] and continues its countdown. It does not wait for an acknowledgment, since the priority of a refresh request supersedes that of any pending non-CPU access and it will be serviced immediately. In this way, a refresh request is guaranteed to occur every (RefreshPeriod[8:0]+1) cycles. A given refresh request may incur some incidental delay in being serviced, due to alignment with DRAM accesses and the possibility of a higher-priority CPU pre-access.

Refresh is also included in the unused read and write timeslot allocation, having second option on awards to a round-robin position shared with the CPU. A refresh issued as a result of an unused timeslot allocation also causes the refresh counter to reload with the value in RefreshPeriod[8:0]. The first access issued by the DAU after reset must be a refresh. This assures that refreshes for all DRAM words fall within the required 3.2 ms window.

//issue a refresh request if counter reaches 0 or at reset or for re-allocated slot if RefreshPeriod != 0 AND (refresh_cnt == 0 OR diu_soft_reset_n == 0 OR prst_n ==0 OR unused_timeslot_allocation == 1) then refresh_req = 1 //de-assert refresh request when refresh acked else if refresh_ack == 1 then refresh_req = 0 //refresh counter if refresh_cnt == 0 OR diu_soft_reset_n == 0 OR prst_n ==0 OR unused timeslot allocation == 1 then refresh_cnt = RefreshPeriod else refresh_cnt = refresh_cnt − 1

Refresh can preceded by a CPU access in the same way as any other access. This is controlled by the CPUPreAccessTimeslots and CPUTotalTimeslots configuration registers. Refresh will therefore not affect CPU performance. A sequence of accesses including refresh might therefore be CPU, refresh, CPU, actual timeslot.

20.14.12.10 CPU Timeslot Controller Description

CPU accesses have priority over all other accesses. CPU access is not included in the timeslot allocations. CPU access is controlled by the CPUPreAccessTimeslots and CPUTotalTimeslots configuration registers.

To avoid the CPU having to wait for its next timeslot it is desirable to have a mechanism for ensuring that the CPU always gets the next available timeslot without incurring any latency on the non-CPU timeslots.

This is be done by defining each timeslot as consisting of a CPU access preceding a non-CPU access. Two counters of 4-bits each are defined allowing the CPU to get a maximum of (CPUPreAccessTimeslots+1) pre-accesses out of a total of (CPUTotalTimeslots+1) main slots. A timeslot counter starts at CPUTotalTimeslots and decrements every timeslot, while another counter starts at CPUPreAccessTimeslots and decrements every timeslot in which the CPU uses its access. If the pre-access entitlement is used up before (CPUTotalTimeslots+1) slots, no further CPU accesses are allowed. When the CPUTotalTimeslots counter reaches zero both counters are reset to their respective initial values.

When CPUPreAccessTimeslots is set to zero then only one pre-access will occur during every (CPUTotalTimeslots+1) slots.

20.14.12.10.1 Conserving CPU Pre-Accesses

In section 20.10.6.2.1 on page 249, it is described how the CPU can be allowed participate in the unused read round-robin scheme. When enabled by the configuration bit EnableCPURoundRobin, the CPU shares a joint position in the round robin with refresh. In this case, the CPU has priority, ahead of refresh, in availing of any unused slot awarded to this position.

Such CPU round-robin accesses do not count towards depleting the CPU's quota of pre-accesses, specified by CPUPreAccessTimeslots. Note that in order to conserve these pre-accesses, the arbitration logic, when faced with the choice of servicing a CPU request either by a pre-access or by an immediately following unused read slot which the CPU is poised to win, will opt for the latter.

20.14.13 Read and Write Data Multiplexor Sub-Block

TABLE 138 Read and Write Multiplexor Sub-block IO Definition Port name Pins I/O Description Clocks and Resets Pclk 1 In System Clock prst_n 1 In System reset, synchronous active low DIU Read Interface to SoPEC Units diu_data 64 Out Data from DIU to SoPEC Units except CPU. First 64-bits is bits 63:0 of 256 bit word Second 64-bits is bits 127:64 of 256 bit word Third 64-bits is bits 191:128 of 256 bit word Fourth 64-bits is bits 255:192 of 256 bit word dram_cpu_(—) 256 Out 256-bit data from DRAM to CPU. data diu_<unit>_(—) 1 Out Signal from DIU telling SoPEC Unit rvalid that valid read data is on the diu_data bus DIU Write Interface to SoPEC Units <unit>_diu_(—) 64 In Data from SoPEC Unit to DIU except data CPU. First 64-bits is bits 63:0 of 256 bit word Second 64-bits is bits 127:64 of 256 bit word Third 64-bits is bits 191:128 of 256 bit word Fourth 64-bits is bits 255:192 of 256 bit word cpu_diu_(—) 128 In Write data from CPU to DIU. wdatat <unit>_diu_(—) 1 In Signal from SoPEC Unit indicating wvalid that data on <unit>_diu_data is valid. Note that “unit” refers to non-CPU requesters only. cpu_diu_(—) 1 In Write enable for the CPU posted wdatavalid write buffer. Also confirms the validity of cpu_diu_wdata. diu_cpu_(—) 1 Out Indicator that the CPU posted write write_rdy buffer is empty. Inputs from CPU Configuration and Arbitration Logic Sub-block arb_gnt 1 In Signal lasting 1 cycle which indi- cates arbitration has occurred and arb_sel is valid. arb_sel 5 In Signal indicating which requesting SoPEC Unit has won arbitration. Encoding is described in Table . dir_sel 2 In Signal indicating which sense of access associated with arb_sel 00: issue non-CPU write 01: read winner 10: write winner 11: refresh winner Outputs to Command Multiplexor Sub-block write_data_(—) 2 Out Signal indicating that valid write valid data is available for the current command. 00 = not valid 01 = CPU write data valid 10 = non-CPU write data valid 11 = both CPU and non-CPU write data valid wdata 256 Out 256-bit non-CPU write data cpu_wdata 32 Out 32-bit CPU write data Inputs from Command Multiplexor Sub-block write_data_(—) 2 In Signal indicating the Command accept Multiplexor has accepted the write data from the write multiplexor 00 = not valid 01 = accepts CPU write data 10 = accepts non-CPU write data 11 = not valid Inputs from DCU dcu_dau_(—) 256 In 256-bit read data from DCU. rdata dcu_dau_(—) 1 In Signal indicating valid read data rvalid on dcu_dau_rdata. Outputs to CPU Configuration and Arbitration Logic Sub-block read_cmd_(—) 2 Out Signal indicating that read multi- rdy plexor is ready for next read read command. 00 = not ready 01 = ready for CPU read 10 = ready for non-CPU read 11 = ready for both CPU and non-CPU reads write_cmd_(—) 2 Out Signal indicating that write rdy multiplexor is ready for next write command. 00 = not ready 01 = ready for CPU write 10 = ready for non-CPU write 11 = ready for both CPU and non-CPU writes Debug Outputs to CPU Configuration and Arbitration Logic Sub-block read_sel 5 Out Signal indicating the SoPEC Unit for which the current read trans- action is occurring. Encoding is described in Table . read_(—) 1 Out Signal indicating that read trans- complete action to SoPEC Unit indicated by read_sel is complete. 20.14.13.1 Read Multiplexor Logic Description

The Read Multiplexor has 2 read channels

-   -   a separate read bus for the CPU, dram_cpu_data[255:0].     -   and a shared read bus for the rest of SoPEC, diu_data[63:0].

The validity of data on the data busses is indicated by signals diu_<unit>_rvalid.

Timing waveforms for non-CPU and CPU DIU read accesses are shown in FIG. 90 and FIG. 91, respectively.

The Read Multiplexor timing is shown in FIG. 127. FIG. 127 shows both CPU and non-CPU reads. Both CPU and non-CPU channels are independent i.e. data can be output on the CPU read bus while non-CPU data is being transmitted in 4 cycles over the shared 64-bit read bus. CPU read data, dram_cpu_data[255.0], is available in the same cycle as output from the DCU. CPU read data needs to be registered immediately on entering the CPU by a flip-flop enabled by the diu_cpu_rvalid signal.

To ease timing, non-CPU read data from the DCU is first registered in the Read Multiplexor by capturing it in the shared read data buffer of FIG. 126 enabled by the dcu_dau_rvalid signal.

The data is then partitioned in 64-bit words on diu_data[63:0].

20.14.13.1.1 Non-CPU Read Data Coherency

Note that for data coherency reasons, a non-CPU read will always result in read data being returned to the requester which includes the after-effects of any pending (i.e. pre-arbitrated, but not yet executed) non-CPU write to the same address, which is currently cached in the non-CPU write buffer. This is shown graphically in Figure n page 319 on page Err r! B kmark not defined.

Should the pending write be partially masked, then the read data returned must take account of that mask. Pending, masked writes by the CDU and SCB, as well as all unmasked non-CPU writes are fully supported.

Since CPU writes are dealt with on a dedicated write channel, no attempt is made to implement coherency between posted, unexecuted CPU writes and non-CPU reads to the same address.

20.14.13.1.2 Read Multiplexor Command Queue

When the Arbitration Logic sub-block issues a read command the associated value of arb_sel[4:0], which indicates which SoPEC Unit has won arbitration, is written into a buffer, the read command queue.

write_en = arb_gnt AND dir_sel[1:0]==“01” if write_en==1 then WRITE arb_sel into read command queue

The encoding of arb_sel[4:0] is given in Table. dir_sel[1:0]==“01” indicates that the operation is a read. The read command queue is shown in FIG. 128.

The command queue could contain values of arb_sel[4:0] for 3 reads at a time.

-   -   In the scenario of FIG. 127 the command queue can contain 2         values of arb_sel[4:0] i.e. for the simultaneous CDU and CPU         accesses.     -   In the scenario of FIG. 130, the command queue can contain 3         values of arb_sel[4:0] i.e. at the time of the second         dcu_dau_rvalid pulse the command queue will contain an         arb_sel[4:0] for the arbitration performed in that cycle, and         the two previous arb_sel[4:0] values associated with the data         for the first two dcu_dau_rvalid pulses, the data associated         with the first dcu_dau_rvalid pulse not having been fully         transfered over the shared read data bus.

The read command queue is specified as 4 deep so it is never expected to fill.

The top of the command queue is a signal read_type[4:0] which indicates the destination of the current read data. The encoding of read_type[4:0] is given in Table.

20.14.13.1.3 CPU Reads

Read data for the CPU goes straight out on dram_cpu_data[255:0] and dcu_dau_rvalid is output on diu_cpu_rvalid.

cpu_read_complete(0) is asserted when a CPU read at the top of the read command queue occurs. cpu_read_complete(0) causes the read command queue to be popped.

cpu_read_complete(0) = (read_type[4:0] == CPU read) AND (dcu_dau_rvalid == 1)

If the current read command queue location points to a non-CPU access and the second read command queue location points to a CPU access then the next dcu_dau_rvalid pulse received is associated with a CPU access. This is the scenario illustrated in FIG. 127. The dcu_dau_rvalid pulse from the DCU must be output to the CPU as diu_cpu_rvalid. This is achieved by using cpu_read_complete(1) to multiplex dcu_dau_rvalid to diu_cpu_rvalid. cpu_read_complete(1) is also used to pop the second from top read command queue location from the read command queue.

cpu_read_complete(1) = (read_type == non-CPU read) AND SECOND (read_type == CPU read) AND (dcu_dau_rvalid == 1) 20.14.13.1.4 Multiplexing dcu_dau_rvalid

read_type[4:0] and cpu_read_complete(1) multiplexes the data valid signal, dcu_dau_rvalID, from the DCU, between the CPU and the shared read bus logic. diu_cpu_rvalid is the read valid signal going to the CPU. noncpu_rvalid is the read valid signal used by the Read Multiplexor control logic to generate read valid signals for non-CPU reads.

if read_type[4:0] == CPU-read then //select CPU diu_cpu_rvalid:= 1 noncpu_rvalid:= 0 if (read_type[4:0]== non-CPU-read) AND SECOND(read_type[4:0]== CPU-read) AND dcu_dau_rvalid == 1 then //select CPU diu_cpu_rvalid:= 1 noncpu_rval_id:= 0 else //select shared read bus logic diu_cpu_rvalid:= 0 noncpu_rvalid:= 1 20.14.13.1.5 Non-CPU Reads

Read data for the shared read bus is registered in the shared read data buffer using noncpu_rvalid. The shared read buffer has 5 locations of 64 bits with separate read pointer, read_ptr[2:0], and write pointer, write_ptr[2:0].

if noncpu_rvalid == 1 and (4 spaces in shared read buffer) then shared_read_data_buffer[write_ptr] = dcu_dau_data[63:0] shared_read_data_buffer[write_ptr+1] = dcu_dau_data[127:64] shared_read_data_buffer[write_ptr+2] = dcu_dau_data[191:128] shared_read_data_buffer[write_ptr+3] = dcu_dau_data[255:192]

The data written into the shared read buffer must be output to the correct SoPEC DIU read requestor according to the value of read_type[4:0] at the top of the command queue. The data is output 64 bits at a time on diu_data[63:0] according to a multiplexor controlled by read_ptr[2:0].

-   -   diu_data[63:0=shared_read_data_buffer[read_ptr]

FIG. 126 shows how read_type[4:0] also selects which shared read bus requesters diu_<unit>_rvalid signal is connected to shared_rvalid. Since the data from the DCU is registered in the Read Multiplexor then shared_rvalid is a delayed version of noncpu_rvalid.

When the read valID, diu_<unit>_rvalID, for the command associated with read_type[4:0] has been asserted for 4 cycles then a signal shared_read_complete is asserted. This indicates that the read has completed. shared_read_complete causes the value of read_type[4:0] in the read command queue to be popped.

A state machine for shared read bus access is shown in FIG. 129. This show the generation of shared_rvalID, shared_read_complete and the shared read data buffer read pointer, read_ptr[2:0], being incremented.

Some points to note from FIG. 129 are:

-   -   shared_rvalid is asserted the cycle after dcu_dau_rvalid         associated with a shared read bus access. This matches the cycle         delay in capturing dau_dcu_data[255:0] in the shared read data         buffer. shared_rvalid remains asserted in the case of back to         back shared read bus accesses.     -   shared_read_complete is asserted in the last shared_rvalid cycle         of a non-CPU access. shared_read_complete causes the shared read         data queue to be popped.         20.14.13.1.6 Read Command Queue Read Pointer Logic

The read command queue read pointer logic works as follows.

if shared_read_complete == 1 OR cpu_read_complete(0) == 1 then POP top of read command queue if cpu_read_complete(1) == 1 then POP second read command queue location 20.14.13.1.7 Debug Signals

shared_read_complete and cpu_read_complete together define read_which indicates to the debug logic that a read has completed. The source of the read is indicated on read_sel[4:0].

read_complete = shared_read_complete OR cpu_read_complete(0) OR cpu_read_complete(1) if cpu_read_complete(1) == 1 then read_sel:= SECOND(read_type) else read_sel:= read_type 20.14.13.1.8 Flow Control

There are separate indications that the Read Multiplexor is able to accept CPU and shared read bus commands from the Arbitration Logic. These are indicated by read_cmd_rdy[1:0].

The Arbitration Logic can always issue CPU reads except if the read command queue fills. The read command queue should be large enough that this should never occur.

//Read Multiplexor ready for Arbitration Logic to issue CPU reads read_cmd_rdy[0] == read command queue not full

For the shared read data, the Read Multiplexor deasserts the shared read bus read_cmd_rdy[1] indication until a space is available in the read command queue. The read command queue should be large enough that this should never occur.

read_cmd_rdy[1] is also deasserted to provide flow control back to the Arbitration Logic to keep the shared read data bus just full.

//Read Multiplexor not ready for Arbitration Logic to issue non-CPU reads read_cmd_rdy[1] = (read command queue not full) AND (flow_control = 0)

The flow control condition is that DCU read data from the second of two back-to-back shared read bus accesses becomes available. This causes read_cmd rdy[1] to de-assert for 1 cycle, resulting in a repeated MSN2 DCU state. The timing is shown in FIG. 130.

flow_control = (read_type[4:0] == non-CPU read) AND SECOND(read_type[4:0] == non- CPU read) AND (current DCU state == MSN2) AND (previous DCU state == MSN1).

FIG. 130 shows a series of back to back transfers over the shared read data bus. The exact timing of the implementation must not introduce any additional latency on shared read bus read transfers i.e. arbitration must be re-enabled just in time to keep back to back shared read bus data full.

The following sequence of events is illustrated in FIG. 130:

-   -   Data from the first DRAM access is written into the shared read         data buffer.     -   Data from the second access is available 3 cycles later, but its         transfer into the shared read buffer is delayed by a cycle, due         to the MSN2 stall condition. (During this delay, read data for         access 2 is maintained at the output of the DRAM.) A similar         1-cycle delay is introduced for every subsequent read access         until the back-to-back sequence comes to an end.     -   Note that arbitration always occurs during the last MSN2 state         of any access. So, for the second and later of any back-to-back         non-CPU reads, arbitration is delayed by one cycle, i.e. it         occurs every fourth cycle instead of the standard every third.

This mechanism provides flow control back to the Arbitration Logic sub-block. Using this mechanism means that the access rate will be limited to which ever takes longer—DRAM access or transfer of read data over the shared read data bus. CPU reads are always be accepted by the Read Multiplexor.

20.14.13.2 Write Multiplexor Logic Description

The Write Multiplexor supplies write data to the DCU.

There are two separate write channels, one for CPU data on cpu_diu_(—)[127:0], one for non-CPU data on non_cpu_wdata[255:0]. A signal write_data valid[1:0] indicates to the Command Multiplexor that the data is valid. The Command Multiplexor then asserts a signal write_data_accept[1:0] indicating that the data has been captured by the DRAM and the appropriate channel in the Write Multiplexor can accept the next write data.

Timing waveforms for write accesses are shown in FIG. 92 to FIG. 94, respectively.

There are 3 types of write accesses:

-   -   CPU accesses

CPU write data on cpu_diu_wdata[127:0] is output on cpu_wdata[127:0]. Since CPU writes are posted, a local buffer is used to store the write data, address and mask until the CPU wins arbitration. This buffer is one position deep. write_data_valid[0], which is synonymous with !diu_cpu_w_write_rdy, remains asserted until the Command Multiplexor indicates it has been written to the DRAM by asserting write_data_accept[0]. The CPU write buffer can then accept new posted writes.

For non-CPU writes, the Write Multiplexor multiplexes the write data from the DIU write requester to the write data buffer and the <unit>_diu_wvalid signal to the write multiplexor control logic.

-   -   CDU accesses         -   64-bits of write data each for a masked write to a separate             256-bit word are transferred to the Write Multiplexor over 4             cycles.         -   When a CDU write is selected the first 64-bits of write data             on cdu_diu_wdata[63:0] are multiplexed to             non_cpu_wdata[63:0]. write_data_valid[1] is asserted to             indicate a non-CPU access when cdu_diu_wvalid is asserted.             The data is also written into the first location in the             write data buffer. This is so that the data can continue to             be output on non_cpu_wdata[63:0] and write_data_valid[1]             remains asserted until the Command Multiplexor indicates it             has been written to the DRAM by asserting             write_data_accept[1]. Data continues to be accepted from the             CDU and is written into the other locations in the write             data buffer. Successive write_data_accept[1] pulses cause             the successive 64-bit data words to be output on wdata[63:0]             together with write_data_valid[1]. The last             write_data_accept[1] means the write buffer is empty and new             write data can be accepted.     -   Other write accesses.         -   256-bits of write data are transferred to the Write             Multiplexor over 4 successive cycles. When a write is             selected the first 64-bits of write data on             <unit>_diu_wdata[63:0] are written into the write data             buffer. The next 64-bits of data are written to the buffer             in successive cycles. Once the last 64-bit word is available             on <unit>_diu_wdata[63:0] the entire word is output on             non_cpu_wdata[255:0], write_data_valid [1] is asserted to             indicate a non-CPU access, and the last 64-bit word is             written into the last location in the write data buffer.             Data continues to be output on non_cpu_wdata[255:0] and             write_data_valid[1] remains asserted until the Command             Multiplexor indicates it has been written to the DRAM by             asserting write_data_accept[1]. New write data can then be             written into the write buffer.

CPU write multiplexor control logic

When the Command Multiplexor has issued the CPU write it asserts write_data_accept[0]. write_data_accept[0] causes the write multiplexor to assert write_cmd_rdy[0].

The signal write_cmd_rdy[0] tells the Arbitration Logic sub-block that it can issue another CPU write command i.e. the CPU write data buffer is empty.

Non-CPU Write Multiplexor Control Logic

The signal write_cmd_rdy[1] tells the Arbitration Logic sub-block that the Write Multiplexor is ready to accept another non-CPU write command. When write_cmd_rdy[1] is asserted the Arbitration Logic can issue a write command to the Write Multiplexor. It does this by writing the value of arb_sel[4:0] which indicates which SoPEC Unit has won arbitration into a write command register, write_cmd[3:0].

write_en = arb_gnt AND dir_sel[1]==1 AND arb_sel = non- CPU if write_en==1 then write_cmd = arb_sel

The encoding of arb_sel[4:0] is given in Table. dir_sel[1]==1 indicates that the operation is a write. arb_sel[4:0] is only written to the write command register if the write is a non-CPU write.

A rule was introduced in Section 20.7.2.3 Interleaving read and write accesses to the effect that non-CPU write accesses would not be allocated adjacent timeslots. This means that a single write command register is required.

The write command register, write_cmd[3:0], indicates the source of the write data. write_cmd[3:0] multiplexes the write data <unit>_diu_wdata, and the data valid signal, <unit>_diu_wvalID, from the selected write requestor to the write data buffer. Note, that CPU write data is not included in the multiplex as the CPU has its own write channel. The <unit>_diu_wvalid are counted to generate the signal word_sel[1:0] which decides which 64-bit word of the write data buffer to store the data from <unit>_diu_wdata.

//when the Command Multiplexor accepts the write data if write_data_accept[1] = 1 then //reset the word select signal word_sel[1:0]=00 //when wvalid is asserted if wvalid = 1 then //increment the word select signal if word_sel[1:0] == 11 then word_sel[1:0] == 00 else word_sel[1:0] == word_sel[1:0] + 1 wvalid is the <unit>_diu_wvalid signal multiplexed by write_cmd[3:0]. word_sel[1:0] is reset when the Command Multiplexor accepts the write data. This is to ensure that word_sel[1:0] is always starts at 00 for the first wvalid pulse of a 4 cycle write data transfer.

The write command register is able to accept the next write when the Command Multiplexor accepts the write data by asserting write_data_accept[1]. Only the last write_data_accept[1] pulse associated with a CDU access (there are 4) will cause the write command register to be ready to accept the next write data.

Flow control back to the Command Multiplexor

-   write_cmd rdy[0] is asserted when the CPU data buffer is empty. -   write_cmd_rdy[1] is asserted when both the write command register     and the write data buffer is empty.     PEP Subsystem     21 PEP Controller Unit (PCU)     21.1 Overview

The PCU has three functions:

-   -   The first is to act as a bus bridge between the CPU-bus and the         PCU-bus for reading and writing PEP configuration registers.     -   The second is to support page banding by allowing the PEP blocks         to be reprogrammed between bands by retrieving commands from         DRAM instead of being programmed directly by the CPU.     -   The third is to send register debug information to the RDU,         within the CPU subsystem, when the PCU is in Debug Mode.         21.2 Interfaces Between PCU and Other Units         21.3 Bus Bridge

The PCU is a bus-bridge between the CPU-bus and the PCU-bus. The PCU is a slave on the CPU-bus but is the only master on the PCU-bus. See Figure page 39 on page Error! Bookmark not defined.

21.3.1 CPU Accessing PEP

All the blocks in the PEP can be addressed by the CPU via the PCU. The MMU in the CPU-subsystem will decode a PCU select signal, cpu_pcu_sel, for all the PCU mapped addresses (see section 11.4.3 on page 69). Using cpu_adr bits 15-12 the PCU will decode individual block selects for each of the blocks within the PEP. The PEP blocks then decode the remaining address bits needed to address their PCU-bus mapped registers. Note: the CPU is only permitted to perform supervisor-mode data-type accesses of the PEP, i.e. cpu_acode=11. If the PCU is selected by the CPU and any other code is present on the cpu_acode bus the access is ignored by the PCU and the pcu_cpu_berr signal is strobed,

CPU commands have priority over DRAM commands. When the PCU is executing each set of four commands retrieved from DRAM the CPU can access PCU-bus registers. In the case that DRAM commands are being executed and the CPU resets the CmdSource to zero, the contents of the DRAM CmdFifo is invalidated and no further commands from the fifo are executed. The CmdPending and NextBandCmdEnable work registers are also cleared.

When a DRAM command writes to the CmdAdr register it means the next DRAM access will occur at the address written to CmdAdr. Therefore if the JUMP instruction is the first command in a group of four, the other three commands get executed and then the PCU will issue a read request to DRAM at the address specified by the JUMP instruction. If the JUMP instruction is the second command then the following two commands will be executed before the PCU requests from the new DRAM address specified by the JUMP instruction etc. Therefore the PCU will always execute the remaining commands in each four command group before carrying out the JUMP instruction.

21.4 Page Banding

The PCU can be programmed to associate microcode in DRAM with each finishedband signal. When a finishedband signal is asserted the PCU will read commands from DRAM and execute these commands. These commands are each 64-bits (see Section 21.8.5) and consist of 32-bit address bits and 32 data bits and allow PCU mapped registers to be programmed directly by the PCU.

If more than one finishedband signal is received at the same time, or others are received while microcode is already executing, the PCU will hold the commands as pending, and will execute them at the first opportunity.

Each microcode program associated with cdu_finishedband, lbd_finishedband and te_finishedband would simply restart the appropriate unit with new addresses—a total of about 4 or 5 microcode instructions. As well, or alternatively, pcu_finishedband can be used to set up all of the units and therefore involves many more instructions. This minimizes the time that a unit is idle in between bands. The pcu_finishedband control signal is issued once the specified combination of CDU, LBD and TE (programmed in BandSelectMask) have finished their processing for a band.

21.5 Interrupts, Address Legality and Security

Interrupts are generated when the various page expansion units have finished a particular band of data from DRAM. The cdu_finishedband, lbd_finishedband and te_finishedband signals are combined in the PCU into a single interrupt pcu_finishedband which is exported by the PCU to the interrupt controller.

The PCU mapped registers should only be accessible from Supervisor Data Mode. The area of DRAM where PCU commands are stored should be a Supervisor Mode only DRAM area, although this is not enforced by the PCU.

When the PCU is executing commands from DRAM, any block-address decoded from a command which is not part of the PEP block-address map will cause the PCU to ignore the command and strobe the pcu_icu_address_invalid interrupt signal. The CPU can then interrogate the PCU to find the source of the illegal command. The MMU will ensure that the CPU cannot address an invalid PEP subsystem block.

When the PCU is executing commands from DRAM, any address decoded from a command which is not part of the PEP address map will cause the PCU to:

-   -   Cease execution of current command and flush all remaining         commands already retrieved from DRAM.     -   Clear CmdPending work-register.     -   Clear NextBandCmdEnable registers.     -   Set CmdSource to zero.

In addition to cancelling all current and pending DRAM accesses the PCU strobes the pcu_icu_address_invalid interrupt signal. The CPU can then interrogate the PCU to find the source of the illegal command.

21.6 Debug Mode

When the need to monitor the (possibly changing) value in any PEP configuration register the PCU may be placed in Debug Mode. This is done via the CPU setting certain Debug Address register within the PCU. Once in Debug Mode the PCU continually reads the target PEP configuration register and sends the read value to the RDU. Debug Mode has the lowest priority of all PCU functions: if the CPU wishes to perform an access or there are DRAM commands to be executed they will interrupt the Debug access, and the PCU will resume Debug access once a CPU or DRAM command has completed.

21.7 Implementation

21.7.1 Definitions of I/O

TABLE 139 PCU Port List Port Name Pins I/O Description Clocks and Resets Pclk 1 In SoPEC functional clock prst_n 1 In Active-low, synchronous reset in pclk domain End of Band Functionality cdu_(—) 1 In Finished band signal from CDU finishedband lbd_(—) 1 In Finished band signal from LBD finishedband te_(—) 1 In Finished band signal from TE finishedband pcu_(—) 1 Out Asserted once the specified finishedband combination of CDU, LBD, and TE have finished their pro- cessing for a band. PCU address error pcu_icu_(—) 1 Out Strobed if PCU decodes a non address_(—) PEP address from commands invalid retrieved from DRAM or CPU. CPU Subsystem Interface Signals cpu_adr[15:2] 14 In CPU address bus. 14 bits are required to decode the address space for the PEP. cpu_dataout 32 In Shared write data bus from [31:0] the CPU pcu_cpu_(—) 32 Out Read data bus to the CPU data[31:0] cpu_rwn 1 In Common read/not-write signal from the CPU cpu_acode 2 In CPU Access Code signals. [1:0] These decode as follows: 00 - User program access 01 - User data access 10 - Supervisor program access 11 - Supervisor data access cpu_pcu_sel 1 In Block select from the CPU. When cpu_pcu_sel is high both cpu_adr and cpu_dataout are valid pcu_cpu_rdy 1 Out Ready signal to the CPU. When pcu_cpu_rdy is high it indi- cates the last cycle of the access. For a write cycle this means cpu_dataout has been registered by the block and for a read cycle this means the data on pcu_cpu_data is valid. pcu_cpu_berr 1 Out Bus error signal to the CPU indicating an invalid access. pcu_cpu_(—) 1 Out Debug Data valid on debug_valid pcu_cpu_data bus. Active high. PCU Interface to PEP blocks pcu_adr[11:2] 10 Out PCU address bus. The 10 least significant bits of cpu_adr [15:2] allow 1024 32-bit word addressable locations per PEP block. Only the number of bits required to decode the address space are exported to each block. pcu_dataout 32 Out Shared write data bus from the [31:0] PCU <unit>_pcu_(—) 32 In Read data bus from each PEP datain[31:0] subblock to the PCU pcu_rwn 1 Out Common read/not-write signal from the PCU pcu_<unit>_(—) 1 Out Block select for each PEP block sel from the PCU. Decoded from the 4 most significant bits of cpu_adr[15:2]. When pcu_<unit>_sel is high both pcu_adr and pcu_dataout are valid <unit>_pcu_(—) 1 In Ready from each PEP block signal rdy to the PCU. When <unit>_pcu_rdy is high it indicates the last cycle of the access. For a write cycle this means pcu_dataout has been registered by the block and for a read cycle this means the data on <unit>_pcu_datain is valid. DIU Read Interface signals pcu_diu_rreq 1 Out PCU requests DRAM read. A read request must be accompanied by a valid read address. pcu_diu_radr 17 Out Read address to DIU [21:5] 17 bits wide (256-bit aligned word). diu_pcu_rack 1 In Acknowledge from DIU that read request has been accepted and new read address can be placed on pcu_diu_radr diu_data[63:0] 64 In Data from DIU to PCU. First 64-bits is bits 63:0 of 256 bit word Second 64-bits is bits 127:64 of 256 bit word Third 64-bits is bits 191:128 of 256 bit word Fourth 64-bits is bits 255:192 of 256 bit word diu_pcu_(—) 1 In Signal from DIU telling PCU rvalid that valid read data is on the diu_data bus 21.7.2 Configuration Registers

TABLE 140 PCU Configuration Registers Address PCU_base+ register #bits reset description Control registers 0x00 Reset 1 0x1 A write to this register causes a reset of the PCU. This register can be read to indicate the reset state: 0 - reset in progress 1 - reset not in progress 0x04 CmdAdr[21:5] 17 0x00 000 The address of the next set of commands to (256-bit aligned retrieve from DRAM. DRAM address) When this register is written to, either by the CPU or DRAM command, 1 is also written to CmdSource to cause the execution of the commands at the specified address. 0x08 BandSelectMask[2:0] 3 0x0 Selects which input finishedBand flags are to be watched to generate the combined pcu_finishedband signal. Bit0 - lbd_finishedband Bit1 - cdu_finishedband Bit2 - te_finishedband 0x0C, NextBandCmdAdr[3:0][21:5] 4 × 17 0x00 000 The address to transfer to CmdAdr as soon 0x10, (256-bit aligned as possible after the next finishedBand[n] 0x14, DRAM address) signal has been received as long as 0x18 NextBandCmdEnable[n] is set. A write from the PCU to NextBandCmdAdr[n] with a non-zero value also sets NextBandCmdEnable[n]. A write from the PCU to NextBandCmdAdr[n] with a 0 value clears NextBandCmdEnable[n]. 0x1C NextCmdAdr[21:5] 17 0x00 000 The address to transfer to CmdAdr when the CPU pending bit (CmdPending[4]) get serviced. A write from the PCU to NextCmdAdr[n] with a non-zero value also sets CmdPending[4]. A write from the PCU to NextCmdAdr[n] with a 0 value clears CmdPending[4] 0x20 CmdSource 1 0x0 0 - commands are taken from the CPU 1 - commands are taken from the CPU as well as DRAM at CmdAdr. 0x24 DebugSelect[15:2] 14 0x00 00 Debug address select. Indicates the address of the register to report on the pcu_cpu_data bus when it is not otherwise being used, and the PEP bus is not being used Bits [15:12] select the unit (see Table ) Bits [11:2] select the register within the unit Work registers (read only) 0x28 InvalidAddress[21:3] 19 0 DRAM Address of current 64-bit command (64-bit aligned attempting to execute. DRAM) Read only register. 0x2C CmdPending 5 0x00 For each bit n, where n is 0 to 3 0 -no commands pending for NextBandCmdAdr[n] 1 -commands pending for NextBandCmdAdr[n] For bit 4 0 -no commands pending for NextCmdAdr[n] 1 -commands pending for NextCmdAdr[n] Read only register. 0x34 FinishedSoFar 3 0x0 The appropriate bit is set whenever the corre- sponding input finishedBand flag is set and the corresponding bit in the BandSelectMask bit is also set. If all FinishedSoFar bits are set wherever BandSelect bits are also set, all FinishedSoFar bits are cleared and the output pcu_finishedband signal is given. Read only register. 0x38 NextBandCmdEnable 4 0x0 This register can be written to indirectly (i.e. the bits are set or cleared via writes to NextBandCmdAdr[n]) For each bit: 0 - do nothing at the next finishedBand[n] signal. 1 - Execute instructions at NextBandCmdAdr[n] as soon as possible after receipt of the next finishedBand[n] signal. Bit0 - lbd_finishedband Bit1 - cdu_finishedband Bit2 - te_finishedband Bit3 - pcu_finishedband Read only register. 21.8 Detailed Description 21.8.1 PEP Blocks Register Map

All PEP accesses are 32-bit register accesses.

From Table 140 it can be seen that four bits only are necessary to address each of the sub-blocks within the PEP part of SoPEC. Up to 14 bits may be used to address any configurable 32-bit register within PEP. This gives scope for 1024 configurable registers per sub-block. This address will come either from the CPU or from a command stored in DRAM. The bus is assembled as follows:

-   -   adr[15:12]=sub-block address     -   adr[n:2]=32-bit register address within sub-block, only the         number of bits required to decode the registers within each         sub-block are used.

TABLE 141 PEP blocks Register Map Block Select Decode = Block cpu_adr[15:12] PCU 0x0 CDU 0x1 CFU 0x2 LBD 0x3 SFU 0x4 TE 0x5 TFU 0x6 HCU 0x7 DNC 0x8 DWU 0x9 LLU 0xA PHI 0xB Reserved 0xC to 0xF 21.8.2 Internal PCU PEP Protocol

The PCU performs PEP configuration register accesses via a select signal, pcu_<block>_sel. The read/write sense of the access is communicated via the pcu_rwn signal (1=read, 0=write). Write data is clocked out, and read data clocked in upon receipt of the appropriate select-read/write-address combination.

FIG. 133 shows a write operation followed by a read operation. The read operation is shown with wait states while the PEP block returns the read data.

For access to the PEP blocks a simple bus protocol is used. The PCU first determines which particular PEP block is being addressed so that the appropriate block select signal can be generated. During a write access PCU write data is driven out with the address and block select signals in the first cycle of an access. The addressed PEP block responds by asserting its ready signal indicating that it has registered the write data and the access can complete. The write data bus is common to all PEP blocks.

A read access is initiated by driving the address and select signals during the first cycle of an access. The addressed PEP block responds by placing the read data on its bus and asserting its ready signal to indicate to the PCU that the read data is valid. Each block has a separate point-to-point data bus for read accesses to avoid the need for a tri-stateable bus.

Consecutive accesses to a PEP block must be separated by at least a single cycle, during which the select signal must be de-asserted.

21.8.3 PCU DRAM Access Requirements

The PCU can execute register programming commands stored in DRAM. These commands can be executed at the start of a print run to initialize all the registers of PEP. The PCU can also execute instructions at the start of a page, and between bands. In the inter-band time, it is critical to have the PCU operate as fast as possible. Therefore in the inter-page and inter-band time the PCU needs to get low latency access to DRAM.

A typical band change requires on the order of 4 commands to restart each of the CDU, LBD, and TE, followed by a single command to terminate the DRAM command stream. This is on the order of 5 commands per restart component.

The PCU does single 256 bit reads from DRAM. Each PCU command is 64 bits so each 256 bit DRAM read can contain 4 PCU commands. The requested command is read from DRAM together with the next 3 contiguous 64-bits which are cached to avoid unnecessary DRAM reads. Writing zero to CmdSource causes the PCU to flush commands and terminate program access from DRAM for that command stream. The PCU requires a 256-bit buffer to the 4 PCU commands read by each 256-bit DRAM access. When the buffer is empty the PCU can request DRAM access again. Adding a 256-bit double buffer would allow the next set of 4 commands to be fetched from DRAM while the current commands are being executed.

1024 commands of 64 bits requires 8 kB of DRAM storage.

Programs stored in DRAM are referred to as PCU Program Code.

21.8.4 End of Band Unit

The state machine is responsible for watching the various input xx_finishedband signals, setting the FinishedSoFar flags, and outputting the pcu_finishedband flags as specified by the BandSelect register.

Each cycle, the end of band unit performs the following tasks:

pcu_finishedband = (FinishedSoFar[0] == BandSelectMask[0]) AND (FinishedSoFar[1] == BandSelectMask[1]) AND (FinishedSoFar[2] == BandSelectMask[2]) AND (BandSelectMask[0] OR BandSelectMask[1] OR BandSelectMask[2]) if (pcu_finishedband == 1) then FinishedSoFar[0] = 0 FinishedSoFar[1] = 0 FinishedSoFar[2] = 0 else FinishedSoFar[0] = (FinishedSoFar[0] OR lbd_finishedband) AND BandSelectMask[0] FinishedSoFar[1] = (FinishedSoFar[1] OR cdu_finishedband) AND BandSelectMask[1] FinishedSoFar[2] = (FinishedSoFar[2] OR te_finishedband) AND BandSelectMask[2]

Note that it is the responsibility of the microcode at the start of printing a page to ensure that all 3 FinishedSoFar bits are cleared. It is not necessary to clear them between bands since this happens automatically.

If a bit of BandSelectMask is cleared, then the corresponding bit of FinishedSoFar has no impact on the generation of pcu_finishedband.

21.8.5 Executing Commands From DRAM

Registers in PEP can be programmed by means of simple 64-bit commands fetched from DRAM. The format of the commands is given in Table 142. Register locations can have a data value of up to 32 bits. Commands are PEP register write commands only.

TABLE 142 Register write commands in PEP bits bits bits bits command 63-32 31-16 15-2 1-0 Register write data zero 32-bit word zero address

Due attention must be paid to the endianness of the processor. The LEON processor is a big-endian processor (bit 7 is the most significant bit).

21.8.6 General Operation

Upon a Reset condition, CmdSource is cleared (to 0), which means that all commands are initially sourced only from the CPU bus interface. Registers and can then be written to or read from one location at a time via the CPU bus interface.

If CmdSource is 1, commands are sourced from the DRAM at CmdAdr and from the CPU bus.

Writing an address to CmdAdr automatically sets CmdSource to 1, and causes a command stream to be retrieved from DRAM. The PCU will execute commands from the CPU or from the DRAM command stream, giving higher priority to the CPU always.

If CmdSource is 0 the DRAM requester examines the CmdPending bits to determine if a new DRAM command stream is pending. If any of CmdPending bits are set, then the appropriate NextBandCmdAdr or NextCmdAdr is copied to CmdAdr (causing CmdSource to get set to 1) and a new command DRAM stream is retrieved from DRAM and executed by the PCU. If there are multiple pending commands the DRAM requestor will service the lowest number pending bit first. Note that a new DRAM command stream only gets retrieved when the current command stream is empty.

If there are no DRAM commands pending, and no CPU commands the PCU defaults to an idle state. When idle the PCU address bus defaults to the DebugSelect register value (bits 11 to 2 in particular) and the default unit PCU data bus is reflected to the CPU data bus. The default unit is determined by the DebugSelect register bits 15 to 12.

In conjunction with this, upon receipt of a finishedBand[n] signal, NextBandCmdEnable[n] is copied to CmdPending[n] and NextBandCmdEnable[n] is cleared. Note, each of the LBD, CDU, and TE (where present) may be re-programmed individually between bands by appropriately setting NextBandCmdAdr[2-0] respectively. However, execution of inter-band commands may be postponed until all blocks specified in the BandSelectMask register have pulsed their finishedband signal. This may be accomplished by only setting NextBandCmdAdr[3] (indirectly causing NextBandCmdEnable[3] to be set) in which case it is the pcu_finishedband signal which causes NextBandCmdEnable[3] to be copied to CmdPending[3].

To conveniently update multiple registers, for example at the start of printing a page, a series of Write Register commands can be stored in DRAM. When the start address of the first Write Register command is written to the CmdAdr register (via the CPU), the CmdSource register is automatically set to 1 to actually start the execution at CmdAdr. Alternatively the CPU can write to NextCmdAdr causing the CmdPending[4] bit to get set, which will then get serviced by the DRAM requester in the pending bit arbitration order.

The final instruction in the command block stored in DRAM must be a register write of 0 to CmdSource so that no more commands are read from DRAM. Subsequent commands will come from pending programs or can be sent via the CPU bus interface.

21.8.6.1 Debug Mode

Debug mode is implemented by reusing the normal CPU and DRAM access decode logic. When in the Arbitrate state (see state machine A below), the PEP address bus is defaulted to the value in the DebugSelect register. The top bits of the DebugSelect register are used to decode a select to a PEP unit and the remaining bits are reflected on the PEP address bus. The selected units read data bus is reflected on the pcu_cpu_data bus to the RDU in the CPU. The pcu_cpu_debug_valid signal indicates to the RDU that the data on the pcu_cpu_data bus is valid debug data.

Normal CPU and DRAM command access will require the PEP bus, and as such will cause the debug data to be invalid during the access, this is indicated to the RDU by setting pcu_cpu_debug_valid to zero.

The decode logic is:

// Default Debug decode if state == Arbitrate then if (cpu_pcu_sel == 1 AND cpu_acode /= SUPERVISOR_DATA_MODE) then pcu_cpu_debug_valid  = 0 // bus error condition pcu_cpu_data = 0 else <unit> = decode(DebugSelect[15:12]) if (<unit> == PCU ) then pcu_cpu_data = Internal PCU register else pcu_cpu_data = <unit>_pcu_datain[31:0] pcu_adr[11:2] = DebugSelect[11:2] pcu_cpu_debug_valid = 1 AFTER 4 clock cycles else pcu_cpu_debug_valid = 0 21.8.7 State Machines

DRAM command fetching and general command execution is accomplished using two state machines. State machine A evaluates whether a CPU or DRAM command is being executed, and proceeds to execute the command(s). Since the CPU has priority over the DRAM it is permitted to interrupt the execution of a stream of DRAM commands.

Machine B decides which address should be used for DRAM access, fetches commands from DRAM and fills a command fifo which A executes. The reason for separating the two functions is to facilitate the execution of CPU or Debug commands while state machine B is performing DRAM reads and filling the command fifo. In the case where state machine A is ready to execute commands (in its Arbitrate state) and it sees both a full DRAM command fifo and an active cpu_pcu_sel then the DRAM commands are executed last.

21.8.7.1 State Machine A: Arbitration and Execution of Commands

The state-machine enters the Reset state when there is an active strobe on either the reset pin, prst_n, or the PCU's soft-reset register. All registers in the PCU are zeroed, unless otherwise specified, on the next rising clock edge. The PCU self-deasserts the soft reset in the pclk cycle after it has been asserted.

The state changes from Reset to Arbitrate when prst_n==1 and PCU_softreset==1.

The state-machine waits in the Arbitrate state until it detects a request for CPU access to the PEP units (cpu_pcu_sel==1 and cpu_acode==11) or a request to execute DRAM commands CmdSource==1, and DRAM commands are available, CmdFifoFull==1. Note if (cpu_pcu_sel==1 and cpu_acode !=11) the CPU is attempting an illegal access. The PCU ignores this command and strobes the cpu_pcu_berr for one cycle.

While in the Arbitrate state the machine assigns the DebugSelect register to the PCU unit decode logic and the remaining bits to the PEP address bus. When in this state the debug data returned from the selected PEP unit is reflected on the CPU bus (pcu_cpu_data bus) and the pcu_cpu_debug valid=1.

If a CPU access request is detected (cpu_pcu_sel==1 and cpu_acode==11) then the machine proceeds to the CpuAccess state. In the CpuAccess state the cpu address is decoded and used to determine the PEP unit to select. The remaining address bits are passed through to the PEP address bus. The machine remains in the CpuAccess state until a valid ready from the selected PEP unit is received. When received the machine returns to the arbitrate state, and the ready signal to the CPU is pulsed.

// decode the logic pcu_<unit>_sel = decode(cpu_adr[15:12]) pcu_adr[11:2] = cpu_adr[11:2]

The CPU is prevented from generating an invalid PEP unit address (prevented in the MMU) and so CPU accesses cannot generate an invalid address error.

If the state machine detects a request to execute DRAM commands (CmdSource==1), it will wait in the Arbitrate state until commands have been loaded into the command FIFO from DRAM (all controlled by state machine B). When the DRAM commands are available (cmd_fifo_full==1) the state machine will proceed to the DRAMAccess state.

When in the DRAMAccess state the commands are executed from the cmd_(—)_fifo. A command in the cmd_fifo consists of 64-bits (or which the FIFO holds 4). The decoding of the 64-bits to commands is given in Table. For each command the decode is

// DRAM command decode pcu_<unit>_sel = decode( cmd_fifo[cmd_count][15:12] ) pcu_adr[11:2] = cmd_fifo[cmd_count][11:2] pcu_dataout = cmd_fifo[cmd_count][63:32]

When the selected PEP unit returns a ready signal (<unit>_pcu_rdy==1) indicating the command has completed, the state machine will return to the Arbitrate state. If more commands exists (cmd_count !=0) the transition will decrement the command count.

When in the DRAMAccess state, if when decoding the DRAM command address bus (cmd_fifo[cmd_count][15.12]), the address selects a reserved address, the state machine proceeds to the AdrError state, and then back to the Arbitrate state. An address error interrupt will be generated and the DRAM command FIFOs will be cleared.

A CPU access can pre-empt any pending DRAM commands. After each command is completed the state machine returns to the Arbitrate state. If a CPU access is required and DRAM command stream is executing the CPU access always takes priority. If a CPU or DRAM command sets the CmdSource to 0, all subsequent DRAM commands in the command FIFO are cleared. If the CPU sets the CmdSource to 0 the CmdPending and NextBandCmdEnable work registers are also cleared.

21.8.7.2 State Machine B: Fetching DRAM Commands

A system reset (prst_n==0) or a software reset (pcu_softreset_n==0) will cause the state machine to reset to the Reset state. The state machine remains in the Reset until both reset conditions are removed. When removed the machine proceeds to the Waft state.

The state machine waits in the Wait state until it determines that commands are needed from DRAM. Two possible conditions exist that require DRAM access. Either the PCU is processing commands which must be fetched from DRAM (cmd_source==1), and the command FIFO is empty (cmd_fifo_full_(—) _(—)=0), or the cmd_source==0 and the command FIFO is empty and there are some commands pending (cmd_pending !=0). In either of these conditions the machine proceeds to the Ack state and issues a read request to DRAM (pcu_diu_rreq==1), it calculates the address to read from dependent on the transition condition. In the command pending transition condition, the highest priority NextBandCmdAdr (or NextCmdAdr) that is pending is used for the read address (pcu_diu_radr) and is also copied to the CmdAdr register. If multiple pending bits are set the lowest pending bits are serviced first. In the normal PCU processing transition the pcu_diu_radr is the CmdAdr register.

When an acknowledge is received from the DRAM the state machine goes to the FillFifo state. In the FillFifo state the machine waits for the DRAM to respond to the read request and transfer data words. On receipt of the first word of data diu_pcu_rvalid==1, the machine stores the 64-bit data word in the command FIFO (cmd_fifo[3]) and transitions to the Data1, Data2, Data3 states each time waiting for a diu_pcu_rvalid==1 and storing the transferred data word to cmd_fifo[2], cmd_fifo[1] and cmd_fifo[0] respectively.

When the transfer is complete the machine returns to the Wait state, setting the cmd_count to 3, the cmd_fifo_full is set to 1 and the CmdAdr is incremented.

If the CPU sets the CmdSource register low while the PCU is in the middle of a DRAM access, the statemachine returns to the Wait state and the DRAM access is aborted.

21.8.7.3 PCU_ICU_Address_Invalid Interrupt

When the PCU is executing commands from DRAM, addresses decoded from commands which are not PCU mapped addresses (4-bits only) will result in the current command being ignored and the pcu_icu_address_invalid interrupt signal is strobed. When an invalid command occurs all remaining commands already retrieved from DRAM are flushed from the CmdFifo, and the CmdPending, NextBandCmdEnable and CmdSource registers are cleared to zero.

The CPU can then interrogate the PCU to find the source of the illegal DRAM command via the InvalidAddress register.

The CPU is prevented by the MMU from generating an invalid address command.

22 Contone Decoder Unit (CDU)

22.1 Overview

The Contone Decoder Unit (CDU) is responsible for performing the optional decompression of the contone data layer.

The input to the CDU is up to 4 planes of compressed contone data in JPEG interleaved format. This will typically be 3 planes, representing a CMY contone image, or 4 planes representing a CMYK contone image. The CDU must support a page of A4 length (11.7 inches) and Letter width (8.5 inches) at a resolution of 267 ppi in 4 colors and a print speed of 1 side per 2 seconds.

The CDU and the other page expansion units support the notion of page banding. A compressed page is divided into one or more bands, with a number of bands stored in memory. As a band of the page is consumed for printing a new band can be downloaded. The new band may be for the current page or the next page. Band-finish interrupts have been provided to notify the CPU of free buffer space.

The compressed contone data is read from the on-chip DRAM. The output of the CDU is the decompressed contone data, separated into planes. The decompressed contone image is written to a circular buffer in DRAM with an expected minimum size of 12 lines and a configurable maximum. The decompressed contone image is subsequently read a line at a time by the CFU, optionally color converted, scaled up to 1600 ppi and then passed on to the HCU for the next stage in the printing pipeline. The CDU also outputs a cdu_finishedband control flag indicating that the CDU has finished reading a band of compressed contone data in DRAM and that area of DRAM is now free. This flag is used by the PCU and is available as an interrupt to the CPU.

22.2 Storage Requirements for Decompressed Contone Data in DRAM

A single SoPEC must support a page of A4 length (11.7 inches) and Letter width (8.5 inches) at a resolution of 267 ppi in 4 colors and a print speed of 1 side per 2 seconds. The printheads specified in the Bi-lithic Printhead Specification [2] have 13824 nozzles per color to provide full bleed printing for A4 and Letter. At 267 ppi, there are 2304 contone pixels⁹ per line represented by 288 JPEG blocks per color. However each of these blocks actually stores data for 8 lines, since a single JPEG block is 8×8 pixels. The CDU produces contone data for 8 lines in parallel, while the HCU processes data linearly across a line on a line by line basis. The contone data is decoded only once and then buffered in DRAM. This means we require two sets of 8 buffer-lines—one set of 8 buffer lines is being consumed by the CFU while the other set of 8 buffer lines is being generated by the CDU. ⁹Pixels may be 8, 16, 24 or 32 bits depending on the number of color planes (8-bits per color)

The buffer requirement can be reduced by using a 1.5 buffering scheme, where the CDU fills 8 lines while the CFU consumes 4 lines. The buffer space required is a minimum of 12 line stores per color, for a total space of 108 KBytes¹⁰. A circular buffer scheme is employed whereby the CDU may only begin to write a line of JPEG blocks (equals 8 lines of contone data) when there are 8-lines free in the buffer. Once the full 8 lines have been written by the CDU, the CFU may now begin to read them on a line by line basis. ¹⁰ 12 lines×4 colors×2304 bytes (assumes 267 ppi, 4 color, full bleed A4/Letter)

This reduction in buffering comes with the cost of an increased peak bandwidth requirement for the CDU write access to DRAM. The CDU must be able to write the decompressed contone at twice the rate at which the CFU reads the data. To allow for trade-offs to be made between peak bandwidth and amount of storage, the size of the circular buffer is configurable. For example, if the circular buffer is configured to be 16 lines it behaves like a double-buffer scheme where the peak bandwidth requirements of the CDU and CFU are equal. An increase over 16 lines allows the CDU to write ahead of the CFU and provides it with a margin to cope with very poor local compression ratIOs in the image.

SoPEC should also provide support for A3 printing and printing at resolutions above 267 ppi. This increases the storage requirement for the decompressed contone data (buffer) in DRAM. Table 143 gives the storage requirements for the decompressed contone data at some sample contone resolutions for different page sizes. It assumes 4 color planes of contone data and a 1.5 buffering scheme.

TABLE 143 Storage requirements for decompressed contone data (buffer) Contone Storage Page resolution Scale Pixels required size (ppi) factor^(a) per line (kBytes) A4/Letter^(b) 267 6 2304 108^(d) 400 4 3456 162 800 2 6912 324 A3^(c) 267 6 3248 152.25 400 4 4872 228.37 800 2 9744 456.75 ^(a)Required for CFU to convert to final output at 1600 dpi ^(b)Bi-lithic printhead has 13824 nozzles per color providing full bleed printing for A4/Letter ^(c)Bi-lithic printhead has 19488 nozzles per color providing full bleed printing for A3 ^(d)12 lines × 4 colors × 2304 bytes. 22.3 Decompression Performance Requirements

The JPEG decoder core can produce a single color pixel every system clock (pclk) cycle, making it capable of decoding at a peak output rate of 8 bits/cycle. SoPEC processes 1 dot (bi-level in 6 colors) per system clock cycle to achieve a print speed of 1 side per 2 seconds for full. bleed A4/Letter printing. The CFU replicates pixels a scale factor (SF) number of times in both the horizontal and vertical directions to convert the final output to 1600 ppi. Thus the CFU consumes a 4 color pixel (32 bits) every SF x SF cycles. The 1.5 buffering scheme described in section 22.2 on page 327 means that the CDU must write the data at twice this rate. With support for 4 colors at 267 ppi, the decompression output bandwidth requirement is 1.78 bits/cycle¹¹. ¹¹2×((4 colors×8 bits)/(6×6 cycles))=1.78 bits/cycle

The JPEG decoder is fed directly from the main memory via the DRAM interface. The amount of compression determines the input bandwidth requirements for the CDU. As the level of compression increases, the bandwidth decreases, but the quality of the final output image can also decrease. Although the average compression ratio for contone data is expected to be 10:1, the average bandwidth allocated to the CDU allows for a local minimum compression ratio of 5:1 over a single line of JPEG blocks. This equates to a peak input bandwidth requirement of 0.36 bits/cycle for 4 colors at 267 ppi, full bleed A4/Letter printing at 1 side per 2 seconds.

Table 144 gives the decompression output bandwidth requirements for different resolutions of contone data to meet a print speed of 1 side per 2 seconds. Higher resolution requires higher bandwidth and larger storage for decompressed contone data in DRAM. A resolution of 400 ppi contone data in 4 colors requires 4 bits/cycle¹², which is practical using a 1.5 buffering scheme. However, a resolution of 800 ppi would require a double buffering scheme (16 lines) so the CDU only has to match the CFU consumption rate. In this case the decompression output bandwidth requirement is 8 bits/cycle¹³, the limiting factor being the output rate of the JPEG decoder core. ¹²2×((4 colors×8 bits)/(4×4 cycles))=4 bits/cycle¹³(4 colors×8 bits)/(2×2 cycles)=8 bits/cycle

TABLE 144 CDU performance requirements for full bleed A4/Letter printing at 1 side per 2 seconds. Contone Scale Decompression output bandwidth resolution (ppi) factor requirement (bits/cycle)^(a) 267 6 1.78 400 4 4 800 2 8^(b) ^(a)Assumes 4 color pixel contone data and a 12 line buffer. ^(b)Scale factor 2 requires at least a 16 line buffer. 22.4 Data Flow

FIG. 136 shows the general data flow for contone data—compressed contone planes are read from DRAM by the CDU, and the decompressed contone data is written to the 12-line circular buffer in DRAM. The line buffers are subsequently read by the CFU.

The CDU allows the contone data to be passed directly on, which will be the case if the color represented by each color plane in the JPEG image is an available ink. For example, the four colors may be C, M, Y, and K, directly represented by CMYK inks. The four colors may represent gold, metallic green etc. for multi-SoPEC printing with exact colors.

However JPEG produces better compression ratIOs for a given visible quality when luminance and chrominance channels are separated. With CMYK, K can be considered to be luminance, but C, M, and Y each contain luminance information, and so would need to be compressed with appropriate luminance tables. We therefore provide the means by which CMY can be passed to SoPEC as YCrCb. K does not need color conversion. When being JPEG compressed, CMY is typically converted to RGB, then to YCrCb and then finally JPEG compressed. At decompression, the YCrCb data is obtained and written to the decompressed contone store by the CDU. This is read by the CFU where the YCrCb can then be optionally color converted to RGB, and finally back to CMY.

The external RIP provides conversion from RGB to YCrCb, specifically to match the actual hardware implementation of the inverse transform within SoPEC, as per CCIR 601-2 [24] except that Y, Cr and Cb are normalized to occupy all 256 levels of an 8-bit binary encoding.

The CFU provides the translation to either RGB or CMY. RGB is included since it is a necessary step to produce CMY, and some printers increase their color gamut by including RGB inks as well as CMYK.

22.5 Implementation

A block diagram of the CDU is shown in FIG. 137.

All output signals from the CDU (cdu_cfu_wradv8line, cdu_finishedband, cdu_icu_jpegerror, and control signals to the DIU) must always be valid after reset. If the CDU is not currently decoding, cdu_cfu_wradv8line, cdu_finishedband and cdu_icu_jpegerror will always be 0.

The read control unit is responsible for keeping the JPEG decoder's input FIFO full by reading compressed contone bytestream from external DRAM via the DIU, and produces the cdu_finishedband signal. The write control unit accepts the output from the JPEG decoder a half JPEG block (32 bytes) at a time, writes it into a double-buffer, and writes the double buffered decompressed half blocks to DRAM via the DIU, interacting with the CFU in order to share DRAM buffers.

22.5.1 Definitions of I/O

TABLE 145 CDU port list and description Port name Pins I/O Description Clocks and reset Pclk 1 In System clock. Jclk 1 In Gated version of system clock used to clock the JPEG decoder core and logic at the output of the core. Allows for stalling of the JPEG core at a pixel sample boundary. jclk_enable 1 Out Gating signal for jclk. prst_n 1 In System reset, synchronous active low. jrst_n 1 In Reset for jclk domain, synchronous active low. PCU interface pcu_cdu_sel 1 In Block select from the PCU. When pcu_cdu_sel is high both pcu_adr and pcu_dataout are valid. pcu_rwn 1 In Common read/not-write signal from the PCU. pcu_adr[7:2] 6 In PCU address bus. Only 6 bits are required to decode the address space for this block. pcu_dataout 32 In Shared write data bus from the [31:0] PCU. cdu_pcu_rdy 1 Out Ready signal to the PCU. When cdu_pcu_rdy is high it indicates the last cycle of the access. For a write cycle this means pcu_dataout has been registered by the block and for a read cycle this means the data on cdu_pcu_datain is valid. cdu_pcu_(—) 32 Out Read data bus to the PCU. datain[31:0] DIU read interface cdu_diu_rreq 1 Out CDU read request, active high. A read request must be accompanied by a valid read address. diu_cdu_rack 1 In Acknowledge from DIU, active high. Indicates that a read request has been accepted and the new read address can be placed on the address bus, cdu_diu_radr. cdu_diu_radr 17 Out CDU read address. 17 bits wide [21:5] (256-bit aligned word). diu_cdu_rvalid 1 In Read data valid, active high. Indicates that valid read data is now on the read data bus, diu_data. diu_data[63:0] 64 In Read data from DRAM. DIU write interface cdu_diu_wreq 1 Out CDU write request, active high. A write request must be accompanied by a valid write address and valid write data. diu_cdu_wack 1 In Acknowledge from DIU, active high. Indicates that a write request has been accepted and the new write address can be placed on the address bus, cdu_diu_wadr. cdu_diu_wadr 19 Out CDU write address. 19 bits wide [21:3] (64-bit aligned word). cdu_diu_wvalid 1 Out Write data valid, active high. Indicates that valid data is now on the write data bus, cdu_diu_data. cdu_diu_data[63:0] 64 Out Write data bus. CFU interface cfu_cdu_rdadvline 1 In Read line pulse, active high. Indicates that the CFU has finished reading a line of decom- pressed contone data to the circular buffer in DRAM and that line of the buffer is now free. cdu_cfu_(—) 1 Out Indicates if the contone line linestore_rdy store has 1 or more lines available to read by the CFU. TE and LBD interface cdu_start_of_(—) 17 Out Points to the 256-bit word that bandstore[21:5] defines the start of the memory area allocated for page bands. cdu_end_of_(—) 17 Out Points to the 256-bit word that bandstore[21:5] defines the last address of the memory area allocated for page bands. ICU interface cdu_finishedband 1 Out CDU's finishedBand flag, active high. Interrupt to the CPU to indicate that the CDU has finished processing a band of compressed contone data in DRAM and that area of DRAM is now free. This signal goes to both the interrupt con- troller and the PCU. cdu_icu_jpegerror 1 Out Active high interrupt indicating an error has occurred in the JPEG decoding process and decompression has stopped. A reset of the CDU must be performed to clear this interrupt. 22.5.2 Configuration Registers

The configuration registers in the CDU are programmed via the PCU interface. Refer to section 21.8.2 on page 321 for the description of the protocol and timing diagrams for reading and writing registers in the CDU. Note that since addresses in SoPEC are byte aligned and the PCU only supports 32-bit register reads and writes, the lower 2 bits of the PCU address bus are not required to decode the address space for the CDU. When reading a register that is less than 32 bits wide zeros should be returned on the upper unused bit(s) of cdu_pcu_datain.

Since the CDU, LBD and TE all access the page band store, they share two registers that enable sequential memory accesses to the page band stores to be circular in nature. Table 146 lists these two registers.

TABLE 146 Registers shared between the CDU, LBD, and TE Address Register # Value on (CDU_base+) name bits reset description Setup registers (remain constant during the processing of multiple bands) 0x80 StartOfBandStore[21:5] 17 0x0_0000 Points to the 256-bit word that defines the start of the memory area allocated for page bands. Circular address generation wraps to this start address. 0x84 EndOfBandStore[21:5] 17 0x1_3FFF Points to the 256-bit word that defines the last address of the memory area allocated for page bands. If the current read address is from this address, then instead of adding 1 to the current address, the current address will be loaded from the StartOfBandStore register.

The software reset logic should include a circuit to ensure that both the pclk and jclk domains are reset regardless of the state of the jclk_enable when the reset is initiated.

The CDU contains the following additional registers:

TABLE 147 CDU registers Address Value on (CDU_base+) Register name #bits reset Description Control registers 0x00 Reset 1 0x1 A write to this register causes a reset of the CDU. This terminates all internal operations within the CS6150. All configuration data previously loaded into the core except for the tables is deleted. 0x04 Go 1 0x0 Writing 1 to this register starts the CDU. Writing 0 to this register halts the CDU. When Go is deasserted the state- machines go to their idle states but all counters and configuration registers keep their values. When Go is asserted all counters are reset, but configuration registers keep their values (i.e. they don't get reset). NextBandEnable is cleared when Go is asserted. The CFU must be started before the CDU is started. Go must remain low for at least 384 jclk cycles after a hardware reset (prst_n = 0) to allow the JPEG core to complete its memory itnitialisation sequence. This register can be read to determine if the CDU is running (1 - running, 0 - stopped). Setup registers 0x0C NumLinesAvail 7 0x0 The number of image lines of data that there is space available for in the decompressed data buffer in DRAM. If this drops < 8 the CDU will stall. In normal operation this value will start off atNumBuffLines and will be decremented by 8 whenever the CDU writes a line of JPEG blocks (8 lines of data) to DRAM and incremented by 1 whenever the CFU reads a line of data from DRAM. NumLinesAvail can be overwritten by the CPU to prevent the CDU from stalling. 0x10 MaxPlane 2 0x0 Defines the number of contone planes − 1. For example, this will be 0 for K (greyscale printing), 2 for CMY, and 3 for CMYK. 0x14 MaxBlock 13 0x000 Number of JPEG MCUs (or JPEG block equivalents, i.e. 8 × 8 bytes) in a line − 1. 0x18 BuffStartAdr[21:7] 15 0x0000 Points to the start of the decompressed contone circular buffer in DRAM, aligned to a half JPEG block boundary. A half JPEG block consists of 4 words of 256-bits, enough to hold 32 contone pixels in 4 colors, i.e. half a JPEG block. 0x1C BuffEndAdr[21:7] 15 0x0000 Points to the start of the last half JPEG block at the end of the decompressed contone circular buffer in DRAM, aligned to a half JPEG block boundary. A half JPEG block consists of 4 words of 256-bits, enough to hold 32 contone pixels in 4 colors, i.e. half a JPEG block. 0x20 NumBuffLines[6:2] 5 0x03 Defines size of buffer in DRAM in terms of the number of decompressed contone lines. The size of the buffer should be a multiple of 4 lines with a minimum size of 8 lines. 0x24 BypassJpg 1 0x0 Determines whether or not the JPEG decoder will be bypassed (and hence pixels are copied directly from input to output) 0 - don't bypass, 1 - bypass Should not be changed between bands. 0x30 NextBandCurrSourceAdr[21:5] 17 0x0_0000 The 256-bit aligned word address containing the start of the next band of compressed contone data in DRAM. This value is copied to CurrSourceAdr when both DoneBand is 1 and NextBandEnable is 1, or when Go transitions from 0 to 1. 0x34 NextBandEndSourceAdr[21:3] 19 0x0_0000 The 64-bit aligned word address containing the last bytes of the next band of compressed contone data in DRAM. This value is copied to EndSourceAdr when when both DoneBand is 1 and NextBandEnable is 1, or when Go transitions from 0 to 1. 0x38 NextBandValid- 3 0x0 Indicates the number of valid bytes − 1 BytesLastFetch in the last 64-bit fetch of the next band of compressed contone data from DRAM. eg 0 implies bits 7:0 are valid, 1 implies bits 15:0 are valid, 7 implies all 63:0 bits are valid etc. This value is copied to ValidBytesLastFetch when both DoneBand is 1 and NextBandEnable is 1, or when Go transitions from 0 to 1. 0x3C NextBandEnable 1 0x0 When NextBandEnable is 1 and DoneBand is 1 NextBandCurrSourceAdr is copied to CurrSourceAdr, NextBandEndSourceAdr is copied to EndSourceAdr NextBandValidBytesLastFetch is copied to ValidBytesLastFetch DoneBand is cleared, NextBandEnable is cleared. NextBandEnable is cleared when Go is asserted. Note that DoneBand gets cleared regardless of the state of Go. Read-only registers 0x40 DoneBand 1 0x0 Specifies whether or not the current band has finished loading into the local FIFO. It is cleared to 0 when Go transitions from 0 to 1. When the last of the compressed contone data for the band has been loaded into the local FIFO, the cdu_finishedband signal is given out and the DoneBandflag is set. If NextBandEnable is 1 at this time then CurrSourceAdr, EndSourceAdr and ValidBytesLastFetch are updated with the values for the next band and DoneBand is cleared. Processing of the next band starts immediately. If NextBandEnable is 0 then the remainder of the CDU will continue to run, decompressing the data already loaded, while the read control unit waits for NextBandEnable to be set before it restarts. 0x44 CurrSourceAdr[21:5] 17 0x0_0000 The current 256-bit aligned word address within the current band of compressed contone data in DRAM. 0x48 EndSourceAdr[21:3] 19 0x0_0000 The 64-bit aligned word address containing the last bytes of the current band of compressed contone data in DRAM. 0x4C ValidBytesLastFetch 3 0x00 Indicates the number of valid bytes − 1 in the last 64-bit fetch of the current band of compressed contone data from DRAM. eg 0 implies bits 7:0 are valid, 1 implies bits 15:0 are valid, 7 implies all 63:0 bits are valid etc. JPEG decoder core setup registers 0x50 JpgDecMask 5 0x00 As segments are decoded they can also be output on the DecJpg (JpgDecHdr) port with the user selecting the segments for output by setting bits in the jpgDecMask port as follows: 4 SOF + SOS + DNL 3 COM + APP 2 DRI 1 DQT 0 DHT If any one of the bits of jpgDecMask is asserted then the SOI and EOI markers are also passed to the DecJpg port. 0x54 JpgDecTType 1 0x0 Test type selector: 0 - DCT coefficients displayed on JpgDecTdata 1 - QDCT coefficient displayed on JpgDecTdata 0x58 JpgDecTestEn 1 0x0 Signal which causes the memories to be bypassed for test purposes. 0x5C JpgDecPType 4 0x0 Signal specifying parameters to be placed on port JpgDecPValue (See Table ). JPEG decoder core read-only status registers 0x60 JpgDecHdr 8 0x00 Selected header segments from the JPEG stream that is currently being decoded. Segments selected using JpgMask. 0x64 JpgDecTData 13 0x0000 12 - TSOS output of CS1650, indicates the first output byte of the first 8 × 8 block of the test data. 11 - TSOB output of CS1650, indicates the first output byte of each 8 × 8 block of test data. 10-0 - 11-bit output test data port - displays DCT coefficients or quantized coefficients depending on value of JpgDecTType. 0x68 JpgDecPValue 16 0x0000 Decoding parameter bus which enables various parameters used by the core to be read. The data available on the PValue port is for information only, and does not contain control signals for the decoder core. 0x6C JpgDecStatus 24 0x00_0000 Bit 23 - jpg_core_stall (if set, indicates that the JPEG core is stalled by gating of jclk as the output JPEG halfblock double-buffers of the CDU are full) Bit 22 - pix_out_valid (This signal is an output from the JPEG decoder core and is asserted when a pixel is being output Bits 21-16 - fifo_contents (Number of bytes in compressed contone FIFO at the input of CDU which feeds the JPEG decoder core) Bits 15-0 are JPEG decoder status outputs from the CS6150 (see Table for description of bits). 22.5.3 Typical Operation

The CDU should only be started after the CFU has been started.

For the first band of data, users set up NextBandCurrSourceAdr, NextBandEndSourceAdr, NextBandValidBytesLastFetch, and the various MaxPlane, MaxBlock, BuffStartBlockAdr, BuffEndBlockAdr and NumBuffLines. Users then set the CDU's Go bit to start processing of the band. When the compressed contone data for the band has finished being read in, the cdu_finishedband interrupt will be sent to the PCU and CPU indicating that the memory associated with the first band is now free. Processing can now start on the next band of contone data.

In order to process the next band NextBandCurrSourceAdr, NextBandEndSourceAdr and NextBandValidBytesLastFetch need to be updated before finally writing a 1 to NextBandEnable. There are 4 mechanisms for restarting the CDU between bands:

-   a. cdu_finishedband causes an interrupt to the CPU. The CDU will     have set its DoneBand bit. The CPU reprograms the     NextBandCurrSourceAdr, NextBandEndSourceAdr and     NextBandValidBytesLastFetch registers, and sets NextBandEnable to     restart the CDU. -   b. The CPU programs the CDU's NextBandCurrSourceAdr,     NextBandCurrEndAdr and NextBandValidBytesLastFetch registers and     sets the NextBandEnable bit before the end of the current band. At     the end of the current band the CDU sets DoneBand. As NextBandEnable     is already 1, the CDU starts processing the next band immediately. -   c. The PCU is programmed so that cdu_finishedband triggers the PCU     to execute commands from DRAM to reprogram the     NextBandCurrSourceAdr, NextBandEndSourceAdr and     NextBandValidBytesLastFetch registers and set the NextBandEnable bit     to start the CDU processing the next band. The advantage of this     scheme is that the CPU could process band headers in advance and     store the band commands in DRAM ready for execution. -   d. This is a combination of b and c above. The PCU (rather than the     CPU in b) programs the CDU's NextBandCurrSourceAdr,     NextBandCurrEndAdr and NextBandValidBytesLastFetch registers and     sets the NextBandEnable bit before the end of the current band. At     the end of the current band the CDU sets DoneBand and pulses     cdu_finishedband. As NextBandEnable is already 1, the CDU starts     processing the next band immediately. Simultaneously,     cdu_finishedband triggers the PCU to fetch commands from DRAM. The     CDU will have restarted by the time the PCU has fetched commands     from DRAM. The PCU commands program the CDU's next band shadow     registers and sets the NextBandEnable bit.

If an error occurs in the JPEG stream, the JPEG decoder will suspend its operation, an error bit will be set in the JpgDecStatus register and the core will ignore any input data and await a reset before starting decoding again. An interrupt is sent to the CPU by asserting cdu_icu_jpegerror and the CDU should then be reset by means of a write to its Reset register before a new page can be printed.

22.5.4 Read Control Unit

The read control unit is responsible for reading the compressed contone data and passing it to the JPEG decoder via the FIFO. The compressed contone data is read from DRAM in single 256-bit accesses, receiving the data from the DIU over 4 clock cycles (64-bits per cycle). The protocol and timing for read accesses to DRAM is described in section 20.9.1 on page 240. Read accesses to DRAM are implemented by means of the state machine described in FIG. 138. All counters and flags should be cleared after reset. When Go transitions from 0 to 1 all counters and flags should take their initial value. While the Go bit is set, the state machine relies on the DoneBand bit to tell it whether to attempt to read a band of compressed contone data. When DoneBand is set, the state machine does nothing. When DoneBand is clear, the state machine continues to load data into the JPEG input FIFO up to 256-bits at a time while there is space available in the FIFO. Note that the state machine has no knowledge about numbers of blocks or numbers of color planes—it merely keeps the JPEG input FIFO full by consecutive reads from DRAM. The DIU is responsible for ensuring that DRAM requests are satisfied at least at the peak DRAM read bandwidth of 0.36 bits/cycle (see section 22.3 on page 329).

A modulo 4 counter, rd_count, is use to count each of the 64-bits received in a 256-bit read access. It is incremented whenever diu_cdu_rvalid is asserted. As each 64-bit value is returned, indicated by diu_cdu_rvalid being asserted, curr_source_adr is compared to both end_source_adr and end_of_bandstore:

-   -   If {curr_source_adr rd_count} equals end_source_adr, the         end_of_band control signal sent to the FIFO is 1 (to signify the         end of the band), the finishedCDUBand signal is output, and the         DoneBand bit is set. The remaining 64-bit values in the burst         from the DIU are ignored, i.e. they are not written into the         FIFO.     -   If rd_count equals 3 and {curr_source_adr, rd_count} does not         equal end_source_adr, then curr_source_adr is updated to be         either start_of_bandstore or curr_source_adr+1, depending on         whether curr_source_adr also equals end_of_bandstore. The         end_of_band control signal sent to the FIFO is 0.

curr_source_adr is output to the DIU as cdu_diu_radr.

A count is kept of the number of 64-bit values in the FIFO. When diu_cdu_rvalid is 1 and ignore_data is 0, data is written to the FIFO by asserting FifoWr, and fifo_contents[3:0] and fifo_wr_adr[2:0] are both incremented.

When fifo_contents[3:0] is greater than 0, jpg_in_strb is asserted to indicate that there is data available in the FIFO for the JPEG decoder core. The JPEG decoder core asserts jpg_in_rdy when it is ready to receive data from the FIFO. Note it is also possible to bypass the JPEG decoder core by setting the BypassJpg register to 1. In this case data is sent directly from the FIFO to the half-block double-buffer. While the JPEG decoder is not stalled (jpg_core_stall equal 0), and jpg_in_rdy (or bypass_jpg) and jpg_in_strb are both 1, a byte of data is consumed by the JPEG decoder core. fifo_rd_adr[5:0] is then incremented to select the next byte. The read address is byte aligned, i.e. the upper 3 bits are input as the read address for the FIFO and the lower 3 bits are used to select a byte from the 64 bits. If fifo_rd_adr[2:0]=111 then the next 64-bit value is read from the FIFO by asserting fifo_rd, and fifo_contents[3:0] is decremented.

22.5.5 Compressed Contone FIFO

The compressed contone FIFO conceptually is a 64-bit input, and 8-bit output FIFO to account for the 64-bit data transfers from the DIU, and the 8-bit requirement of the JPEG decoder.

In reality, the FIFO is actually 8 entries deep and 65-bits wide (to accommodate two 256-bit accesses), with bits 63-0 carrying data, and bit 64 containing a 1-bit end_of_band flag. Whenever 64-bit data is written to the FIFO from the DIU, an end of band flag is also passed in from the read control unit. The end_of_band bit is 1 if this is the last data transfer for the current band, and 0 if it is not the last transfer. When end_of_band=1 during an input, the ValidBytesLastFetch register is also copied to an image version of the same.

On the JPEG decoder side of the FIFO, the read address is byte aligned, i.e. the upper 3 bits are input as the read address for the FIFO and the lower 3 bits are used to select a byte from the 64 bits (1st byte corresponds to bits 7-0, second byte to bits 15-8 etc.). If bit 64 is set on the read, bits 63-0 contain the end of the bytestream for that band, and only the bytes specified by the image of ValidBytesLastFetch are valid bytes to be read and presented to the JPEG decoder. Note that ValidBytesLastFetch is copied to an image register as it may be possible for the CDU to be reprogrammed for the next band before the previous band's compressed contone data has been read from the FIFO (as an additional effect of this, the CDU has a non-problematic limitation in that each band of contone data must be more than 4×64-bits, or 32 bytes, in length).

22.5.6 CS6150 JPEG Decoder

JPEG decoder functionality is implemented by means of a modified version of the Amphion CS6150 JPEG decoder core. The decoder is run at a nominal clock speed of 160 MHz. (Amphion have stated that the CS6150 JPEG decoder core can run at 185 MHz in 0.13 um technology). The core is clocked by jclk which a gated version of the system clock pclk. Gating the clock provides a mechanism for stalling the JPEG decoder on a single color pixel-by-pixel basis. Control of the flow of output data is also provided by the PixOutEnab input to the JPEG decoder. However, this only allows stalling of the output at a JPEG block boundary and is insufficient for SoPEC. Thus gating of the clock is employed and PixOutEnab is instead tied high.

The CS6150 decoder automatically extracts all relevant parameters from the JPEG bytestream and uses them to control the decoding of the image. The JPEG bytestream contains data for the Huffman tables, quantization tables, restart interval definition and frame and scan headers. The decoder parses and checks the JPEG bytestream automatically detecting and processing all the JPEG marker segments. After identifying the JPEG segments the decoder re-directs the data to the appropriate units to be stored or processed as appropriate. Any errors detected in the bytestream, apart from those in the entropy coded segments, are signalled and, if an error is found, the decoder stops reading the JPEG stream and waits to be reset.

JPEG images must have their data stored in interleaved format with no subsampling. Images longer than 65536 lines are allowed: these must have an initial imageHeight of 0. If the image has a Define Number Lines (DNL) marker at the end (normally necessary for standard JPEG, but not necessary for SoPEC's version of the CS6150), it must be equal to the total image height mod 64 k or an error will be generated.

See the CS6150 Databook [21] for more details on how the core is used, and for timing diagrams of the interfaces. Note that [21] does not describe the use of the DNL marker in images of more than 64 k lines length as this is a modification to the core.

The CS6150 decoder can be bypassed by setting the BypassJpg register. If this register is set, then the data read from DRAM must be in the same format as if it was produced by the JPEG decoder: 8×8 blocks of pixels in the correct color order. The data is uncompressed and is therefore lossless.

The following subsections describe the means by which the CS6150 internals can be made visible.

22.5.6.1 JPEG Decoder Reset

The JPEG decoder has 2 possible types of reset, an asynchronous reset and a synchronous clear. In SoPEC the asynchronous reset is connected to the hardware synchronous reset of the CDU and can be activated by any hardware reset to SoPEC (either from external pin or from any of the wake-up sources, e.g. USB activity, Wake-up register timeout) or by resetting the PEP section (ResetSection register in the CPR block).

The synchronous clear is connected to the software reset of the CDU and can be activated by the low to high transition of the Go register, or a software reset via the Reset register.

The 2 types of reset differ, in that the asynchronous reset, resets the JPEG core and causes the core to enter a memory initialization sequence that takes 384 clock cycles to complete after the reset is deasserted. The synchronous clear resets the core, but leaves the memory as is. This has some implications for programming the CDU.

In general the CDU should not be started (i.e. setting Go to 1) until at least 384 cycles after a hardware reset. If the CDU is started before then, the memory initialization sequence will be terminated leaving the JPEG core memory in an unknown state. This is allowed if the memory is to be initialized from the incoming JPEG stream.

22.5.6.2 JPEG Decoder Parameter Bus

The decoding parameter bus JpgDecPValue is a 16-bit port used to output various parameters extracted from the input data stream and currently used by the core. The 4-bit selector input (JpgDecPType) determines which internal parameters are displayed on the parameter bus as per Table 148. The data available on the PValue port does not contain control signals used by the CS6150.

TABLE 148 Parameter bus definitions PType Output orientation PValue 0x0 FY[15:0] FY: number of lines in frame 0x1 FX[15:0] FX: number of columns in frame 0x2 00_YMCU[13:0] YMCU: number of MCUs in Y direction of the current scan 0x3 00_XMCU[13:0] XMCU: number of MCUs in X direction of the current scan 0x4 Cs0[7:0]_Tq0[1:0]_(—) Cs0: identifier for the first V0[2:0]_H0[2:0] scan component Tq0: quantization table identifier for the first scan component V0: vertical sampling factor for the first scan component. Values = 1-4 H0: horizontal sampling factor for the first scan component. Values = 1-4 0x5 Cs1[7:0]_Tq1[1:0]_(—) Cs1, Tq1, V1 and H1 for the V1[2:0]_H1[2:0] second scan component. V1, H1 undefined if NS < 2 0x6 Cs2[7:0]_Tq2[1:0]_(—) Cs2, Tq2, V2 and H2 for the V2[2:0]_H2[2:0] second scan component. V2, H2 undefined if NS < 3 0x7 Cs3[7:0]_Tq3[1:0]_(—) Cs3, Tq3, V3 and H3 for the V3[2:0]_H3[2:0] second scan component. V3, H3 undefined if NS < 4 0x8 CsH[15:0] CsH: no. of rows in current scan 0x9 CsV[15:0] CsV: no. of columns in current scan 0xA DRI[15:0] DRI: restart interval 0xB 000_HMAX[2:0]_(—) HMAX: maximal horizontal sampling VMAX[2:0]_(—) factor in frame VMAX: maximal MCUBLK[3:0]_NS[2:0] vertical sampling factor in frame MCUBLK: number of blocks per MCU of the current scan, from 1 to 10 NS: number of scan components in current scan, 1-4 22.5.6.3 JPEG Decoder Status Register

The status register flags indicate the current state of the CS6150 operation. When an error is detected during the decoding process, the decompression process in the JPEG decoder is suspended and an interrupt is sent to the CPU by asserting cdu_icu_jpegerror (generated from DecError). The CPU can check the source of the error by reading the JpgDecStatus register. The CS6150 waits until a reset process is invoked by asserting the hard reset prst_n or by a soft reset of the CDU. The individual bits of JpgDecStatus are set to zero at reset and active high to indicate an error condition as defined in Table 149.

Note: A DecHfError will not block the input as the core will try to recover and produce the correct amount of pixel data. The DecHfError is cleared automatically at the start of the next image and so no intervention is required from the user. If any of the other errors occur in the decode mode then, following the error cancellation, the core will discard all input data until the next Start Of Image (SOI) without triggering any more errors.

The progress of the decoding can be monitored by observing the values of TbIDef, IDctInProg, DecInProg and JpgInProg.

TABLE 149 JPEG decoder status register definitions Bit Name Description 15-12 TblDef[7:4] Indicates the number of Huffman tables defined, 1 bit/table. 11-8  TblDef[3:0] Indicates the number of quantization tables defined, 1 bit/table. 7 DecHfError Set when an undefined Huffman table symbol is referenced during decoding. 6 CtlError Set when an invalid SOF parameter or an invalid SOS parameter is detected. Also set when there is a mismatch between the DNL segment input to the core and the number of lines in the input image which have already been decoded. Note that SoPEC's implementation of the CS6150 does not require a final DNL when the initial setting for ImageHeight is 0. This is to allow images longer than 64k lines. 5 HtError Set when an invalid DHT segment is detected. 4 QtError Set when an invalid DQT segment is detected. 3 DecError Set when anything other than a JPEG marker is input. Set when any of DecFlags[6:4] are set. Set when any data other than the SOI marker is detected at the start of a stream. Set when any SOF marker is detected other than SOF0. Set if incomplete Huffman or quantization definition is detected. 2 IDctInProg Set when IDCT starts processing first data of a scan. Cleared when IDCT has processed the last data of a scan. 1 DecInProg For each scan this signal is asserted after the SigSOS (Start of Scan Segment) signal has been output from the core and is deasserted when the decoding of a scan is complete. It indicates that the core is in the decoding state. 0 JpgInProg Set when core starts to process input data (JpgIn) and de-asserted when decoding has been completed i.e. when the last pixel of last block of the image is output. 22.5.7 Half-block Buffer Interface

Since the CDU writes 256 bits (4×64 bits) to memory at a time, it requires a double-buffer of 2×256 bits at its output. This is implemented in an 8×64 bit FIFO. It is required to be able to stall the JPEG decoder core at its output on a half JPEG block boundary, i.e. after 32 pixels (8 bits per pixel). We provide a mechanism for stalling the JPEG decoder core by gating the clock to the core (with jclk_enable) when the FIFO is full. The output FIFO is responsible for providing two buffered half JPEG blocks to decouple JPEG decoding (read control unit) from writing those JPEG blocks to DRAM (write control unit). Data coming in is in 8-bit quantities but data going out is in 64-bit quantities for a single color plane.

22.5.8 Write Control Unit

A line of JPEG blocks in 4 colors, or 8 lines of decompressed contone data, is stored in DRAM with the memory arrangement as shown FIG. 139. The arrangement is in order to optimize access for reads by writing the data so that 4 color components are stored together in each 256-bit DRAM word.

The CDU writes 8 lines of data in parallel but stores the first 4 lines and second 4 lines separately in DRAM. The write sequence for a single line of JPEG 8×8 blocks in 4 colors, as shown in FIG. 139, is as follows below and corresponds to the order in which pixels are output from the JPEG decoder core:

block 0, color 0, line 0 in word p bits 63–0, line 1 in word p+1 bits 63–0, line 2 in word p+2 bits 63–0, line 3 in word p+3 bits 63–0, block 0, color 0, line 4 in word q bits 63–0, line 5 in word q+1 bits 63–0, line 6 in word q+2 bits 63–0, line 7 in word q+3 bits 63–0, block 0, color 1, line 0 in word p bits 127–64, line 1 in word p+1 bits 127–64, line 2 in word p+2 bits 127–64, line 3 in word p+3 bits 127–64, block 0, color 1, line 4 in word q bits 127–64, line 5 in word q+1 bits 127–64, line 6 in word q+2 bits 127–64, line 7 in word q+3 bits 127–64, repeat for block 0 color 2, block 0 color 3........ block 1, color 0, line 0 in word p+4 bits 63–0, line 1 in word p+5 bits 63–0, etc................................................... block N, color 3, line 4 in word q+4n bits 255–192, line 5 in word q+4n+1 bits 255–192, line 6 in word q+4n+2 bits 255– 192, line 7 in word q+4n+3 bit 255–192

In SoPEC data is written to DRAM 256 bits at a time. The DIU receives a 64-bit aligned address from the CDU, i.e. the lower 2 bits indicate which 64-bits within a 256-bit location are being written to. With that address the DIU also receives half a JPEG block (4 lines) in a single color, 4×64 bits over 4 cycles. All accesses to DRAM must be padded to 256 bits or the bits which should not be written are masked using the individual bit write inputs of the DRAM. When writing decompressed contone data from the CDU, only 64 bits out of the 256-bit access to DRAM are valID, and the remaining bits of the write are masked by the DIU. This means that the decompressed contone data is written to DRAM in 4 back-to-back 64-bit write masked accesses to 4 consecutive 256-bit DRAM locations/words.

Writing of decompressed contone data to DRAM is implemented by the state machine in FIG. 140. The CDU writes the decompressed contone data to DRAM half a JPEG block at a time, 4×64 bits over 4 cycles. All counters and flags should be cleared after reset. When Go transitions from 0 to 1 all counters and flags should take their initial value. While the Go bit is set, the state machine relies on the half_block_ok_to_read and line_store_ok_to_write flags to tell it whether to attempt to write a half JPEG block to DRAM. Once the half-block buffer interface contains a half JPEG block, the state machine requests a write access to DRAM by asserting cdu_diu_wreq and providing the write address, corresponding to the first 64-bit value to be written, on cdu_diu_wadr (only the address the first 64-bit value in each access of 4×64 bits is issued by the CDU. The DIU can generate the addresses for the second, third and fourth 64-bit values). The state machine then waits to receive an acknowledge from the DIU before initiating a read of 4×64-bit values from the half-block buffer interface by asserting rd_adv for 4 cycles. The output cdu_diu_wvalid is asserted in the cycle after rd_adv to indicate to the DIU that valid data is present on the cdu_diu_data bus and should be written to the specified address in DRAM. A rd_adv_half block pulse is then sent to the half-block buffer interface to indicate that the current read buffer has been read and should now be available to be written to again. The state machine then returns to the request state.

The pseudocode below shows how the write address is calculated on a per clock cycle basis.

Note counters and flags should be cleared after reset. When Go transitions from 0 to 1 all counters and flags should be cleared and lwr_halfblock_adr gets loaded with buff_start_adr and upr_halfblock_adr gets loaded with buff_start_adr+max_block+1.

// assign write address output to DRAM cdu_diu_wadr[6:5] = 00 // corresponds to linenumber, only first address is // issued for each DRAM access. Thus line is always 0. // The DIU generates these bits of the address. cdu_diu_wadr[4:3] = color if (half == 1) then cdu_diu_wadr[21:7] = upr_halfblock_adr // for lines 4–7 of JPEG block else cdu_diu_wadr[21:7] = lwr_halfblock_adr // for lines 0–3 of JPEG block // update half, color, block and addresses after each DRAM write access if (rd_adv_half_block == 1) then if (half == 1) then half = 0 if (color == max_plane) then color = 0 if (block == max_block) then // end of writing a line of JPEG blocks pulse wradv8line block = 0 // update half block address for start of next line of JPEG blocks taking // account of address wrapping in circular buffer and 4 line offset if (upr_halfblock_adr == buff_end_adr) then upr_halfblock_adr = buff_start_adr + max_block + 1 elsif (upr_halfblock_adr + max_block + 1 == buff_end_adr) then upr_halfblock_adr = buff_start_adr else upr_halfblock_adr = upr_halfblock_adr + max_block + 2 else block ++ upr_halfblock_adr ++ // move to address for lines 4–7 for next block else color ++ else half = 1 if (color == max_plane) then if (block == max_block) then // end of writing a line of JPEG blocks // update half block address for start of next line of JPEG blocks taking // account of address wrapping in circular buffer and 4 line offset if (lwr_halfblock_adr == buff_end_adr) then lwr_halfblock_adr = buff_start_adr + max_block + 1 elsif (lwr_halfblock_adr + max_block + 1 == buff_end_adr) then lwr_halfblock_adr = buff_start_adr else lwr_halfblock_adr = lwr_halfblock_adr  + max_block + 2 else lwr_halfblock_adr ++ // move to address for lines 0–3 for next block 22.5.9 Contone Line Store Interface

The contone line store interface is responsible for providing the control over the shared resource in DRAM. The CDU writes 8 lines of data in up to 4 color planes, and the CFU reads them line-at-a-time. The contone line store interface provides the mechanism for keeping track of the number of lines stored in DRAM, and provides signals so that a given line cannot be read from until the complete line has been written.

The CDU writes 8 lines of data in parallel but writes the first 4 lines and second 4 lines to separate areas in DRAM. Thus, when the CFU has read 4 lines from DRAM that area now becomes free for the CDU to write to. Thus the size of the line store in DRAM should be a multiple of 4 lines.

The minimum size of the line store interface is 8 lines, providing a single buffer scheme. Typical sizes are 12 lines for a 1.5 buffer scheme while 16 lines provides a double-buffer scheme.

The size of the contone line store is defined by num_buff_lines. A count is kept of the number of lines stored in DRAM that are available to be written to. When Go transitions from 0 to 1, NumLinesAvail is set to the value of num_buff_lines. The CDU may only begin to write to DRAM as long as there is space available for 8 lines, indicated when the line_store_ok_to_write bit is set.

When the CDU has finished writing 8 lines, the write control unit sends an wradv8line pulse to the contone line store interface, and NumLinesAvail is decremented by 8. The write control unit then waits for line_store_ok_to_write to be set again.

If the contone line store is not empty (has one or more lines available in it), the CDU will indicate to the CFU via the cdu_cfu_linestore_rdy signal. The cdu_cfu_linestore_rdy signal is generated by comparing the NumLinesAvail with the programmed num_buff_lines. As the CFU reads a line from the contone line store it will pulse the rdadvline to indicate that it has read a full line from the line store. NumLinesAvail is incremented by 1 on receiving a rdadvline pulse.

To enable running the CDU while the CFU is not running the NumLinesAvail register can also be updated via the configuration register interface. In this scenario the CPU polls the value of the NumLinesAvail register and overwrites it to prevent stalling of the CDU (NumLinesAvail<8). The CPU will always have priority in any updating of the NumLinesAvail register.

23 Contone FIFO Unit (CFU)

23.1 Overview

The Contone FIFO Unit (CFU) is responsible for reading the decompressed contone data layer from the circular buffer in DRAM, performing optional color conversion from YCrCb to RGB followed by optional color inversion in up to 4 color planes, and then feeding the data on to the HCU. Scaling of data is performed in the horizontal and vertical directions by the CFU so that the output to the HCU matches the printer resolution. Non-integer scaling is supported in both the horizontal and vertical directions. Typically, the scale factor will be the same in both directions but may be programmed to be different.

23.2 Bandwidth Requirements

The CFU must read the contone data from DRAM fast enough to match the rate at which the contone data is consumed by the HCU.

Pixels of contone data are replicated a X scale factor (SF) number of times in the X direction and Y scale factor (SF) number of times in the Y direction to convert the final output to 1600 dpi. Replication in the X direction is performed at the output of the CFU on a pixel-by-pixel basis while replication in the Y direction is performed by the CFU reading each line a number of times, according to the Y-scale factor, from DRAM. The HCU generates 1 dot (bi-level in 6 colors) per system clock cycle to achieve a print speed of 1 side per 2 seconds for full bleed A4/Letter printing. The CFU output buffer needs to be supplied with a 4 color contone pixel (32 bits) every SF cycles. With support for 4 colors at 267 ppi the CFU must read data from DRAM at 5.33 bits/cycle¹⁴. ¹⁴32 bits/6 cycles=5.33 bits/cycle

23.3 Color Space Conversion

The CFU allows the contone data to be passed directly on, which will be the case if the color represented by each color plane in the JPEG image is an available ink. For example, the four colors may be C, M, Y, and K, directly represented by CMYK inks. The four colors may represent gold, metallic green etc. for multi-SoPEC printing with exact colors.

JPEG produces better compression ratIOs for a given visible quality when luminance and chrominance channels are separated. With CMYK, K can be considered to be luminance, but C, M and Y each contain luminance information and so would need to be compressed with appropriate luminance tables. We therefore provide the means by which CMY can be passed to SoPEC as YCrCb. K does not need color conversion.

When being JPEG compressed, CMY is typically converted to RGB, then to YCrCb and then finally JPEG compressed. At decompression, the YCrCb data is obtained, then color converted to RGB, and finally back to CMY.

The external RIP provides conversion from RGB to YCrCb, specifically to match the actual hardware implementation of the inverse transform within SoPEC, as per CCIR 601-2 [24] except that Y, Cr and Cb are normalized to occupy all 256 levels of an 8-bit binary encoding.

The CFU provides the translation to either RGB or CMY. RGB is included since it is a necessary step to produce CMY, and some printers increase their color gamut by including RGB inks as well as CMYK.

Consequently the JPEG stream in the color space convertor is one of:

-   -   1 color plane, no color space conversion     -   2 color planes, no color space conversion     -   3 color planes, no color space conversion     -   3 color planes YCrCb, conversion to RGB     -   4 color planes, no color space conversion     -   4 color planes YCrCbX, conversion of YCrCb to RGB, no color         conversion of X

The YCrCb to RGB conversion is described in [14]. Note that if the data is non-compressed, there is no specific advantage in performing color conversion (although the CDU and CFU do permit it).

23.4 Color Space Inversion

In addition to performing optional color conversion the CFU also provides for optional bit-wise inversion in up to 4 color planes. This provides the means by which the conversion to CMY may be finalised, or to may be used to provide planar correlation of the dither matrices.

The RGB to CMY conversion is given by the relationship:

-   -   C=255−R     -   M=255−G     -   Y=255−B

These relationships require the page RIP to calculate the RGB from CMY as follows:

-   -   R=255−C     -   G=255−M     -   B=255−Y         23.5 Scaling

Scaling of pixel data is performed in the horizontal and vertical directions by the CFU so that the output to the HCU matches the printer resolution. The CFU supports non-integer scaling with the scale factor represented by a numerator and a denominator. Only scaling up of the pixel data is allowed, i.e. the numerator should be greater than or equal to the denominator. For example, to scale up by a factor of two and a half, the numerator is programmed as 5 and the denominator programmed as 2.

Scaling is implemented using a counter as described in the pseudocode below. An advance pulse is generated to move to the next dot (x-scaling) or line (y-scaling).

if (count + denominator − numerator >= 0) then count = count + denominator − numerator advance = 1 else count = count + denominator advance = 0 23.6 Lead-In and Lead-Out Clipping

The JPEG algorithm encodes data on a block by block basis, each block consists of 64 8-bit pixels (representing 8 rows each of 8 pixels). If the image is not a multiple of 8 pixels in X and Y then padding must be present. This padding (extra pixels) will be present after decoding of the JPEG bytestream.

Extra padded lines in the Y direction (which may get scaled up in the CFU) will be ignored in the HCU through the setting of the BottomMargin register.

Extra padded pixels in the X direction must also be removed so that the contone layer is clipped to the target page as necessary.

In the case of a multi-SoPEC system, 2 SoPECs may be responsible for printing the same side of a page, e.g. SoPEC #1 controls printing of the left side of the page and SoPEC #2 controls printing of the right side of the page and shown in FIG. 141. The division of the contone layer between the 2 SoPECs may not fall on a 8 pixel (JPEG block) boundary. The JPEG block on the boundary of the 2 SoPECs (JPEG block n below) will be the last JPEG block in the line printed by SoPEC #1 and the first JPEG block in the line printed by SoPEC #2. Pixels in this JPEG block not destined for SoPEC #1 are ignored by appropriately setting the LeadOutClipNum. Pixels in this JPEG block not destined for SoPEC #2 must be ignored at the beginning of each line. The number of pixels to be ignored at the start of each line is specified by the LeadInClipNum register. It may also be the case that the CDU writes out more JPEG blocks than is required to be read by the CFU, as shown for SBPEC #2 below. In this case_the value of the MaxBlock register in the CDU is set to correspond to JPEG block m but the value for the MaxBlock register in the CFU is set to correspond to JPEG block m-1. Thus JPEG block m is not read in by the CFU. Additional clipping on contone pixels is required when they are scaled up to the printer's resolution. The scaling of the first valid pixel in the line is controlled by setting the XstartCount register. The HcuLineLength register defines the size of the target page for the contone layer at the printer's resolution and controls the scaling of the last valid pixel in a line sent to the HCU.

23.7 Implementation

FIG. 142 shows a block diagram of the CFU.

23.7.1 Definitions of I/O

TABLE 150 CFU port list and description Port Name Pins I/O Description Clocks and reset pclk 1 In System clock prst_n 1 In System reset, synchronous active low. PCU interface pcu_cfu_sel 1 In Block select from the PCU. When pcu_cfu_sel is high both pcu_adr and pcu_dataout are valid. pcu_rwn 1 In Common read/not-write signal from the PCU. pcu_adr[6:2] 4 In PCU address bus. Only 5 bits are required to decode the address space for this block. pcu_dataout 32 In Shared write data bus from [31:0] the PCU. cfu_pcu_rdy 1 Out Ready signal to the PCU. When cfu_pcu_rdy is high it indicates the last cycle of the access. For a write cycle this means pcu_dataout has been registered by the block and for a read cycle this means the data on cfu_pcu_datain is valid. cfu_pcu_(—) 32 Out Read data bus to the PCU. datain[31:0] DIU interface cfu_diu_rreq 1 Out CFU read request, active high. A read request must be accom- panied by a valid read address. diu_cfu_rack 1 In Acknowledge from DIU, active high. Indicates that a read request has been accepted and the new read address can be placed on the address bus, cfu_diu_radr. cfu_diu_(—) 17 Out CFU read address. 17 bits radr[21:5] wide (256-bit aligned word). diu_cfu_rvalid 1 In Read data valid, active high. Indicates that valid read data is now on the read data bus, diu_data. diu_data[63:0] 64 In Read data from DRAM. CDU interface cdu_cfu_(—) 1 In When high indicates that the linestore_rdy contone line store has 1 or more lines available to be read by the CFU. cfu_cdu_(—) 1 Out Read line pulse, active high. rdadvline Indicates that the CFU has finished reading a line of decompressed contone data to the circular buffer in DRAM and that line of the buffer is now free. HCU interface hcu_cfu_advdot 1 In Informs the CFU that the HCU has captured the pixel data on cfu_hcu_c[0-3] data lines and the CFU can now place the next pixel on the data lines. cfu_hcu_avail 1 Out Indicates valid data present on cfu_hcu_c[0-3] data lines. cfu_hcu_c0data 8 Out Pixel of data in contone [7:0] plane 0. cfu_hcu_c1data 8 Out Pixel of data in contone [7:0] plane 1. cfu_hcu_c2data 8 Out Pixel of data in contone [7:0] plane 2. cfu_hcu_c3data 8 Out Pixel of data in contone [7:0] plane 3. 23.7.2 Configuration Registers

The configuration registers in the CFU are programmed via the PCU interface. Refer to section 21.8.2 on page 321 for the description of the protocol and timing diagrams for reading and writing registers in the CFU. Note that since addresses in SoPEC are byte aligned and the PCU only supports 32-bit register reads and writes, the lower 2 bits of the PCU address bus are not required to decode the address space for the CFU. When reading a register that is less than 32 bits wide zeros should be returned on the upper unused bit(s) of cfu_pcu_datain. The configuration registers of the CFU are listed in Table 151:

TABLE 151 CFU registers Value Address on (CFU_base+) Register Name #bits Reset Description Control registers 0x00 Reset 1 0x1 A write to this register causes a reset of the CFU. 0x04 Go 1 0x0 Writing 1 to this register starts the CFU. Writing 0 to this register halts the CFU. When Go is deasserted the state- machines go to their idle states but all counters and configuration registers keep their values. When Go is asserted all counters are reset, but configuration registers keep their values (i.e. they don't get reset). The CFU must be started before the CDU is started. This register can be read to determine if the CFU is running (1 - running, 0 - stopped). Setup registers 0x10 MaxBlock 13 0x000 Number of JPEG MCUs (or JPEG block equivalents, i.e. 8 × 8 bytes) in a line − 1. 0x14 BuffStartAdr[21:7] 15 0x0000 Points to the start of the decompressed contone circular buffer in DRAM, aligned to a half JPEG block boundary. A half JPEG block consists of 4 words of 256-bits, enough to hold 32 contone pixels in 4 colors, i.e. half a JPEG block 0x18 BuffEndAdr[21:7] 15 0x0000 Points to the end of the decompressed contone circular buffer in DRAM, aligned to a half JPEG block boundary (address is inclusive). A half JPEG block consists of 4 words of 256-bits, enough to hold 32 contone pixels in 4 colors, i.e. half a JPEG block. 0x1C 4LineOffset 13 0x0000 Defines the offset between the start of one 4 line store to the start of the next 4 line store − 1. In Figure n page 394 on page Error! Bookmark not defined., if BufStartAdr corresponds to line 0 block 0 then BuffStartAdr + 4LineOffset corresponds to line 4 block 0. 4LineOffset is specified in units of 128 bytes, eg 0-128 bytes, 1-256 bytes etc. This register is required in addition to MaxBlock as the number of JPEG blocks in a line required by the CFU may be different from the number of JPEG blocks in a line written by the CDU. 0x20 YCrCb2RGB 1 0x0 Set this bit to enable conversion from YCrCb to RGB. Should not be changed between bands. 0x24 InvertColorPlane 4 0x0 Set these bits to perform bit-wise inversion on a per color plane basis. bit0 - 1 invert color plane 0 - 0 do not convert bit1 - 1 invert color plane 1 - 0 do not convert bit2 - 1 invert color plane 2 - 0 do not convert bit3 - 1 invert color plane 3 Should not be changed between bands. 0x28 HcuLineLength 16 0x0000 Number of contone pixels − 1 in a line (after scaling). Equals the number of hcu_cfu_dotadv pulses − 1 received from the HCU for each line of contone data. 0x2C LeadInClipNum 3 0x0 Number of contone pixels to be ignored at the start of a line (from JPEG block 0 in a line). They are not passed to the output buffer to be scaled in the X direction. 0x30 LeadOutClipNum 3 0x0 Number of contone pixels to be ignored at the end of a line (from JPEG block MaxBlock in a line). They are not passed to the output buffer to be scaled in the X direction. 0x34 XstartCount 8 0x00 Value to be loaded at the start of every line into the counter used for scaling in the X direction. Used to control the scaling of the first pixel in a line to be sent to the HCU. This value will typically be zero, except in the case where a number of dots are clipped on the lead in to a line. 0x38 XscaleNum 8 0x01 Numerator of contone scale factor in X direction. 0x3C XscaleDenom 8 0x01 Denominator of contone scale factor in X direction. 0x40 YscaleNum 8 0x01 Numerator of contone scale factor in Y direction. 0x44 YscaleDenom 8 0x01 Denominator of contone scale factor in Y direction. 23.7.3 Storage of Decompressed Contone Data in DRAM

The CFU reads decompressed contone data from DRAM in single 256-bit accesses. JPEG blocks of decompressed contone data are stored in DRAM with the memory arrangement as shown The arrangement is in order to optimize access for reads by writing the data so that 4 color components are stored together in each 256-bit DRAM word. The means that the CFU reads 64-bits in 4 colors from a single line in each 256-bit DRAM access.

The CFU reads data line at a time in 4 colors from DRAM. The read sequence, as shown in FIG. 143, is as follows:

line 0, block 0 in word p of DRAM line 0, block 1 in word p+4 of DRAM ......................................... line 0, block n in word p+4n of DRAM (repeat to read line a number of times according to scale factor) line 1, block 0 in word p+1 of DRAM line 1, block 1 in word p+5 of DRAM etc......................................

The CFU reads a complete line in up to 4 colors a Y scale factor number of times from DRAM before it moves on to read the next. When the CFU has finished reading 4 lines of contone data that 4 line store becomes available for the CDU to write to.

23.7.4 Decompressed Contone Buffer

Since the CFU reads 256 bits (4 colors x 64 bits) from memory at a time, it requires storage of at least 2×256 bits at its input. To allow for all possible DIU stall conditions the input buffer is increased to 3×256 bits to meet the CFU target bandwidth requirements. The CFU receives the data from the DIU over 4 clock cycles (64-bits of a single color per cycle). It is implemented as 4 buffers. Each buffer conceptually is a 64-bit input and 8-bit output buffer to account for the 64-bit data transfers from the DIU, and the 8-bit output per color plane to the color space converter.

On the DRAM side, wr_buff indicates the current buffer within each triple-buffer that writes are to occur to. wr_sel selects which triple-buffer to write the 64 bits of data to when wr_en is asserted. On the color space converter side, rd_buff indicates the current buffer within each triple-buffer that reads are to occur from. When rd_en is asserted a byte is read from each of the triple-buffers in parallel. rd_sel is used to select a byte from the 64 bits (1st byte corresponds to bits 7-0, second byte to bits 15-8 etc.).

Due to the limitations of available register arrays in IBM technology, the decompressed contone buffer is implemented as a quadruple buffer. While this offers some benefits for the CFU it is not necessitated by the bandwidth requirements of the CFU.

23.7.5 Y-scaling Control Unit

The Y-scaling control unit is responsible for reading the decompressed contone data and passing it to the color space converter via the decompressed contone buffer. The decompressed contone data is read from DRAM in single 256-bit accesses, receiving the data from the DIU over 4 clock cycles (64-bits per cycle). The protocol and timing for read accesses to DRAM is described in section 20.9.1 on page 240. Read accesses to DRAM are implemented by means of the state machine described in FIG. 144.

All counters and flags should be cleared after reset. When Go transitions from 0 to 1 all counters and flags should take their initial value. While the Go bit is set, the state machine relies on the line8_ok_to_read and buff_ok_to_write flags to tell it whether to attempt to read a line of compressed contone data from DRAM. When line8_ok_to_read is 0 the state machine does nothing. When line8_ok_to_read is 1 the state machine continues to load data into the decompressed contone buffer up to 256-bits at a time while there is space available in the buffer. A bit is kept for the status of each 64-bit buffer: buff_avail[0] and buff_avail[1]. It also keeps a single bit (rd_buff) for the current buffer that reads are to occur from, and a single bit (wr_buff) for the current buffer that writes are to occur to.

-   -   buff ok_to_write equals ˜buff_avail[wr_buff]. When a wr_adv_buff         pulse is received, buff_avail[wr_buff] is set, and wr_buff is         inverted. Whenever diu_cfu_rvalid is asserted, wr_en is asserted         to write the 64-bits of data from DRAM to the buffer selected by         wr_sel and wr_buff.     -   buff ok_to_read equals buff_avail[rd_buff]. If there is data         available in the buffer and the output double-buffer has space         available (outbuff_ok_to_write equals 1) then data is read from         the buffer by asserting rd_en and rd_sel gets incremented to         point to the next value. wr_adv is asserted in the following         cycle to write the data to the output double-buffer of the CFU.         When finished reading the buffer, rd_sel equals b111 and rd_en         is asserted, buff_avail[rd_buff] is set, and rd_buff is         inverted.

Each line is read a number of times from DRAM, according to the Y-scale factor, before the CFU moves on to start reading the next line of decompressed contone data. Scaling to the printhead resolution in the Y direction is thus performed.

The pseudocode below shows how the read address from DRAM is calculated on a per clock cycle basis. Note all counters and flags should be cleared after reset or when Go is cleared. When a 1 is written to Go, both curr_halfblock and line_start_halfblock get loaded with buff_start_adr, and y_scale_count gets loaded with y_scale_denom. Scaling in the Y direction is implemented by line replication by re-reading lines from DRAM. The algorithm for non-integer scaling is described in the pseudocode below.

// assign read address output to DRAM cdu_diu_wadr[21:7] = curr_halfblock cdu_diu_wadr[6:5]   = line[1:0] // update block, line, y_scale_count and addresses after each DRAM read access if (wr_adv_buff == 1) then if (block == max_block) then // end of reading a line of contone in up to 4 colors block = 0 // check whether to advance to next line of contone data in DRAM if (y_scale_count + y_scale_denom − y_scale_num >= 0) then y_scale_count = y_scale_count + y_scale_denom − y_scale_num pulse RdAdvline if (line == 3) then // end of reading 4 line store of contone data line = 0 // update half block address for start of next line taking account of // address wrapping in circular buffer and 4 line offset if (curr_halfblock == buff_end_adr) then curr_halfblock = buff_start_adr line_start_adr = buff_start_adr elsif ((line_start_adr + 4line_offset) == buff_end_adr)) then curr_halfblock = buff_start_adr line_start_adr = buff_start_adr else curr_halfblock = line_start_adr + 4line_offset line_start_adr = line_start_adr + 4line_offset else line ++ curr_halfblock = line_start_adr else // re-read current line from DRAM y_scale_count = y_scale_count + y_scale_denom curr_halfblock = line_start_adr else block ++ curr_halfblock ++ 23.7.6 Contone Line Store Interface

The contone line store interface is responsible for providing the control over the shared resource in DRAM. The CDU writes 8 lines of data in up to 4 color planes, and the CFU reads them line-at-a-time. The contone line store interface provides the mechanism for keeping track of the number of lines stored in DRAM, and provides signals so that a given line cannot be read from until the complete line has been written.

A count is kept of the number of lines that have been written to DRAM by the CDU and are available to be read by the CFU. At start-up, buff_lines_avail is set to the 0. The CFU may only begin to read from DRAM when the CDU has written 8 complete lines of contone data. When the CDU has finished writing 8 lines, it sends an cdu_cfu_wradv8line pulse to the CFU, and buff_lines_avail is incremented by 8. The CFU may continue reading from DRAM as long as buff_lines_avail is greater than 0. line8_ok_to_read is set while buff_lines_avail is greater than 0. When it has completely finished reading a line of contone data from DRAM, the Y-scaling control unit sends a RdAdvLine signal to contone line store interface and to the CDU to free up the line in the buffer in DRAM. buff_lines_avail is decremented by 1 on receiving a RdAdvline pulse.

23.7.7 Color Space Converter (CSC)

The color space converter consists of 2 stages: optional color conversion from YCrCb to RGB followed by optional bit-wise inversion in up to 4 color planes.

The convert YCrCb to RGB block takes 3 8-bit inputs defined as Y, Cr, and Cb and outputs either the same data YCrCb or RGB. The YCrCb2RGB parameter is set to enable the conversion step from YCrCb to RGB. If YCrCb2RGB equals 0, the conversion does not take place, and the input pixels are passed to the second stage. The 4th color plane, if present, bypasses the convert YCrCb to RGB block. Note that the latency of the convert YCrCb to RGB block is 1 cycle. This latency should be equalized for the 4th color plane as it bypasses the block.

The second stage involves optional bit-wise inversion on a per color plane basis under the control of invert_color_plane. For example if the input is YCrCbK, then YCrCb2RGB can be set to 1 to convert YCrCb to RGB, and invert_color_plane can be set to 0111 to then convert the RGB to CMY, leaving K unchanged.

If YCrCb2RGB equals 0 and invert_color_plane equals 0000, no color conversion or color inversion will take place, so the output pixels will be the same as the input pixels.

FIG. 145 shows a block diagram of the color space converter.

The convert YCrCb to RGB block is an implementation of [14]. Although only 10 bits of coefficients are used (1 sign bit, 1 integer bit, 8 fractional bits), full internal accuracy is maintained with 18 bits. The conversion is implemented as follows:

-   -   R*=Y+(359/256)(Cr−128)     -   G*=Y−(183/256)(Cr−128)−(88/256)(Cb−128)     -   B*=Y+(454/256)(Cb−128)

R*, G* and B* are rounded to the nearest integer and saturated to the range 0-255 to give R, G and B. Note that, while a Reset results in all-zero output, a zero input gives output RGB=[0¹⁵, 136¹⁶, 0¹⁷]. ///¹⁵⁻179 is saturated to 0¹⁶135.5, with rounding becomes 136.¹⁷⁻227 is saturated to 0

23.7.8 X-scaling Control Unit

The CFU has a 2×32-bit double-buffer at its output between the color space converter and the HCU. The X-scaling control unit performs the scaling of the contone data to the printers output resolution, provides the mechanism for keeping track of the current read and write buffers, and ensures that a buffer cannot be read from until it has been written to.

A bit is kept for the status of each 32-bit buffer: buff_avail[0] and buff_avail[1]. It also keeps a single bit (rd_buff) for the current buffer that reads are to occur from, and a single bit (wr_buff) for the current buffer that writes are to occur to.

The output value outbuff_ok_to_write equals ˜buff_avail[wr_buff]. Contone pixels are counted as they are received from the Y-scaling control unit, i.e. when wr_adv is 1. Pixels in the lead-in and lead-out areas are ignored, i.e. they are not written to the output buffer. Lead-in and lead-out clipping of pixels is implemented by the following pseudocode that generates the wr_en pulse for the output buffer.

if (wradv == 1) then if (pixel_count == {max_block,b111}) then pixel_count = 0 else pixel_count ++ if ((pixel_count < leadin_clip_num) OR (pixel_count > ({max_block,b111} − leadout_clip_num))) then wr_en = 0 else wr_en = 1

When a wr_en pulse is sent to the output double-buffer, buff_avail[wr_buff] is set, and wr_buff is inverted.

The output cfu_hcu_avail equals buff_avail[rd_buff]. When cfu_hcu_avail equals 1, this indicates to the HCU that data is available to be read from the CFU. The HCU responds by asserting hcu_cfu_advdot to indicate that the HCU has captured the pixel data on cfu_hcu_c[0-3]data lines and the CFU can now place the next pixel on the data lines.

The input pixels from the CSC may be scaled a non-integer number of times in the X direction to produce the output pixels for the HCU at the printhead resolution. Scaling is implemented by pixel replication. The algorithm for non-integer scaling is described in the pseudocode below. Note, x_scale_count should be loaded with x_start_count after reset and at the end of each line. This controls the amount by which the first pixel is scaled by. hcu_line_length and hcu_cfu_dotadv control the amount by which the last pixel in a line that is sent to the HCU is scaled by.

if (hcu_cfu_dotadv == 1) then if (x_scale_count + x_scale_denom − x_scale_num >= 0) then x_scale_count = x_scale_count + x_scale_denom − x_scale_num rd_en = 1 else x_scale_count = x_scale_count + x_scale_denom rd_en = 0 else x_scale_count = x_scale_count rd_en = 0

When a rd_en pulse is received, buff_avail[rd_buff] is cleared, and rd_buff is inverted.

A 16-bit counter, dot_adv_count, is used to keep a count of the number of hcu_cfu_dotadv pulses received from the HCU. If the value of dot_adv_count equals hcu_line_length and a hcu_cfu_dotadv pulse is received, then a rd_en pulse is genrated to present the next dot at the output of the CFU, dot_adv_count is reset to 0 and x_scale_count is loaded with x_start_count.

24 Lossless Bi-level Decoder (LBD)

24.1 Overview

The Lossless Bi-level Decoder (LBD) is responsible for decompressing a single plane of bi-level data. In SoPEC bi-level data is limited to a single spot color (typically black for text and line graphics).

The input to the LBD is a single plane of bi-level data, read as a bitstream from DRAM. The LBD is programmed with the start address of the compressed data, the length of the output (decompressed) line, and the number of lines to decompress. Although the requirement for SoPEC is to be able to print text at 10:1 compression, the LBD can cope with any compression ratio if the requested DRAM access is available. A pass-through mode is provided for 1:1 compression. Ten-point plain text compresses with a ratio of about 50:1. Lossless bi-level compression across an average page is about 20:1 with 10:1 possible for pages which compress poorly.

The output of the LBD is a single plane of decompressed bi-level data. The decompressed bi-level data is output to the SFU (Spot FIFO Unit), and in turn becomes an input to the HCU (Halftoner/Compositor unit) for the next stage in the printing pipeline. The LBD also outputs a lbd_finishedband control flag that is used by the PCU and is available as an interrupt to the CPU.

24.2 Main Features OF LBD

FIG. 147 shows a schematic outline of the LBD and SFU.

The LBD is required to support compressed images of up to 800 dpi. If possible we would like to support bi-level images of up to 1600 dpi. The line buffers must therefore be long enough to store a complete line at 1600 dpi.

The PEC1 LBD is required to output 2 dots/cycle to the HCU. This throughput capability is retained for SoPEC to minimise changes to the block, although in SoPEC the HCU will only read 1 dot/cycle. The PEC1 LDB outputs 16 bits in parallel to the PEC1 spot buffer. This is also retained for SoPEC. Therefore the LBD in SoPEC can run much faster than is required. This is useful for allowing stalls, e.g. due to band processing latency, to be absorbed.

The LBD has a pass through mode to cope with local negative compression. Pass through mode is activated by a special run-length code. Pass through mode continues to either end of line or for a pre-programmed number of bits, whichever is shorter. The special run-length code is always executed as a run-length code, followed by pass through.

The LBD outputs decompressed bi-level data to the NextLineFiFO in the Spot FIFO Unit (SFU). This stores the decompressed lines in DRAM, with a typical minimum of 2 lines stored in DRAM, nominally 3 lines up to a programmable number of lines. The SFU's NextLineFIFO can fill while the SFU waits for write access to DRAM. Therefore the LBD must be able to support stalling at its coding process. This is provided by the SFU via it's output during a line.

The LBD uses the previous line in the decoding process. This is provided by the SFU via it's PrevLineFIFO. Decoding can stall in the LBD while this FIFO waits to be filled from DRAM. A signal sfu_lbd_rdy indicates that both the SFU's NextLineFIFO and PrevLineFIFO are available for writing and reading, respectively.

A configuration register in the LBD controls whether the first line being decoded at the start of a band uses the previous line read from the SFU or uses an all 0's line instead.

The line length is stored in DRAM must be programmable to a value greater than 128. An A4 line of 13824 dots requires 1.7 Kbytes of storage. An A3 line of 19488 dots requires 2.4 Kbytes of storage.

The compressed spot data can be read at a rate of 1 bit/cycle for pass through mode 1:1 compression.

The LBD finished band signal is exported to the PCU and is additionally available to the CPU as an interrupt.

24.2.1 Bi-level Decoding in the LBD

The black bi-level layer is losslessly compressed using Silverbrook Modified Group 4 (SMG4) compression which is a version of Group 4 Facsimile compression [22] without Huffman and with simplified run length encodings. The encoding are listed in Table 152 and Table 153.

TABLE 152 Bi-Level group 4 facsimile style compression encodings Encoding Description same as Group 1000 Pass Command: a0

b2, skip next 4 Facsimile two edges 1 Vertical(0): a0

b1, color = !color 110 Vertical(1): a0

b1 + 1, color = !color 010 Vertical(−1): a0

b1 − 1, color = !color 110000 Vertical(2): a0

b1 + 2, color = !color 010000 Vertical(−2): a0

b1 − 2, color = !color Unique to this 100000 Vertical(3): a0

b1 + 3, color = implementation !color 000000 Vertical(−3): a0

b1 − 3, color = !color <RL><RL>100 Horizontal: a0

a0 + <RL> + <RL>

SMG4 has a pass through mode to cope with local negative compression. Pass through mode is activated by a special run-length code. Pass through mode continues to either end of line or for a pre-programmed number of bits, whichever is shorter. The special run-length code is always executed as a run-length code, followed by pass through. The pass through escape code is a medium length run-length with a run of less than or equal to 31.

TABLE 153 Run length (RL) encodings Encoding Description Unique RRRRR1 Short Black Runlength to this (5 bits) implemen- tation RRRRR1 Short White Runlength (5 bits) RRRRRRRRRR10 Medium Black Runlength (10 bits) RRRRRRRR10 Medium White Runlength (8 bits) RRRRRRRRRR10 Medium Black Runlength with RRRRRRRRRR <= 31, Enter pass through RRRRRRRR10 Medium White Runlength with RRRRRRRR <= 31, Enter pass through RRRRRRRRRRRRRRR00 Long Black Runlength (15 bits) RRRRRRRRRRRRRRR00 Long White Runlength (15 bits)

Since the compression is a bitstream, the encodings are read right (least significant bit) to left (most significant bit). The run lengths given as RRRRR in Table 153 are read in the same way (least significant bit at the right to most significant bit at the left).

There is an additional enhancement to the G4 fax algorithm, it relates to pass through mode. It is possible for data to compress negatively using the G4 fax algorithm. On occasions like this it would be easier to pass the data to the LBD as un-compressed data. Pass through mode is a new feature that was not implemented in the PEC1 version of the LBD. When the LBD is in pass through mode the least significant bit of the data stream is an un-compressed bit. This bit is used to construct the current line.

To enter pass through mode the LBD takes advantage of the way run lengths can be written.

Usually if one of the runlength pair is less than or equal to 31 it should be encoded as a short runlength. However under the coding scheme of Table it is still legal to write it as a medium or long runlength. The LBD has been designed so that if a short runlength value is detected in a medium runlength then once the horizontal command containing this runlength is decoded completely this will tell the LBD to enter pass through mode and the bits following the runlength is un-compressed data. The number of bits to pass through is either a programmed number of bits or the end of the line which ever comes first. Once the pass through mode is completed the current color is the same as the color of the last bit of the passed through data.

24.2.2 DRAM Access Requirements

The compressed page store for contone, bi-level and raw tag data is 2 Mbytes. The LBD will access the compressed page store in single 256-bit DRAM reads. The LBD will need a 256-bit double buffer in its interface to the DIU. The LBD's DIU bandwidth requirements are summarized in Table 154

TABLE 154 DRAM bandwidth requirements Maximum number of Peak Average cycles between each Bandwidth Bandwidth Direction 256-bit DRAM access (bits/cycle) (bits/cycle) Read 2561 (1:1 1 (1:1 0.1 (10:1 compression) compression) compression)

1: At 1:1 compression the LBD requires 1 bit/cycle or 256 bits every 256 cycles.

24.3 Implementation

24.3.1 Definitions of IO

TABLE 155 LBD Port List Port Name Pins I/O Description Clocks and Resets Pclk 1 In SoPEC Functional clock. prst_n 1 In Global reset signal. Bandstore signals cdu_endofbandstore[21:5] 17 In Address of the end of the current band of data. 256-bit word aligned DRAM address. cdu_startofbandstore[21:5] 17 In Address of the start of the current band of data. 256-bit word aligned DRAM address. lbd_finishedband 1 Out LBD finished band signal to PCU and Interrupt Controller. DIU Interface signals lbd_diu_rreq 1 Out LBD requests DRAM read. A read request must be accompanied by a valid read address. lbd_diu_radr[21:5] 17 Out Read address to DIU 17 bits wide (256-bit aligned word). diu_lbd_rack 1 In Acknowledge from DIU that read request has been accepted and new read address can be placed on lbd_diu_radr. diu_data[63:0] 64 In Data from DIU to SoPEC Units. First 64-bits is bits 63:0 of 256 bit word. Second 64-bits is bits 127:64 of 256 bit word. Third 64-bits is bits 191:128 of 256 bit word. Fourth 64-bits is bits 255:192 of 256 bit word. diu_lbd_rvalid 1 In Signal from DIU telling SoPEC Unit that valid read data is on the diu_data bus PCU Interface data and control signals pcu_addr[5:2] 4 In PCU address bus. Only 4 bits are required to decode the address space for this block. pcu_dataout[31:0] 32 In Shared write data bus from the PCU. lbd_pcu_datain[31:0] 32 Out Read data bus from the LBD to the PCU. pcu_rwn 1 In Common read/not-write signal from the PCU. pcu_lbd_sel 1 In Block select from the PCU. When pcu_lbd_sel is high both pcu_addr and pcu_dataout are valid. lbd_pcu_rdy 1 Out Ready signal to the PCU. When lbd_pcu_rdy is high it indicates the last cycle of the access. For a write cycle this means pcu_dataout has been registered by the block and for a read cycle this means the data on lbd_pcu_datain is valid. SFU Interface data and control signals sfu_lbd_rdy 1 In Ready signal indicating SFU has previous line data available for reading and is also ready to be written to. lbd_sfu_advline 1 Out Advance line signal to previous and next line buffers lbd_sfu_pladvword 1 Out Advance word signal for previous line buffer. sfu_lbd_pldata[15:0] 16 In Data from the previous line buffer. lbd_sfu_wdata[15:0] 16 Out Write data for next line buffer. lbd_sfu_wdatavalid 1 Out Write data valid signal for next line buffer data. 24.3.2 Configuration Registers

TABLE 156 LBD Configuration Registers Value Address Register on (LBD_base+) Name #Bits Reset Description Control registers 0x00 Reset 1 0x1 A write to this register causes a reset of the LBD. This register can be read to indicate the reset state: 0 - reset in progress 1 - reset not in progress 0x04 Go 1 0x0 Writing 1 to this register starts the LBD. Writing 0 to this register halts the LBD. The Go register is reset to 0 by the LBD when it finishes processing a band. When Go is deasserted the state-machines go to their idle states but all counters and configuration registers keep their values. When Go is asserted all counters are reset, but configuration registers keep their values (i.e. they don't get reset). The LBD should only be started after the SFU is started. This register can be read to determine if the LBD is running (1 - running, 0 - stopped). Setup registers (constant for during processing the page) 0x08 LineLength 16 0x0000 Width of expanded bi-level line (in dots) (must be set greater than 128 bits). 0x0C PassThroughEnable 1 0x1 Writing 1 to this register enables passthrough mode. Writing 0 to this register disables pass- through mode thereby making the LBD compatible with PEC1. 0x10 PassThroughDotLength 16 0x0000 This is the dot length − 1 for which pass- through mode will last. If the end of the line is reached first then pass-through will be disabled. The value written to this register must be a non-zero value. Work registers (need to be set up before processing a band) 0x14 NextBandCurrReadAdr[21:5] 17 0x00000 Shadow register which is copied to (256-bit aligned CurrReadAdr when (NextBandEnable == 1 & DRAM address) Go == 0). NextBandCurrReadAdr is the address of the start of the next band of compressed bi-level data in DRAM. 0x18 NextBandLinesRemaining 15 0x0000 Shadow register which is copied to LinesRemaining when (NextBandEnable == 1 & Go == 0). NextBandLinesRemaining is the number of lines to be decoded in the next band of compressed bi-level data. 0x1C NextBandPrevLineSource 1 0x0 Shadow register which is copied to PrevLineSource when (NextBandEnable == 1 & Go == 0). 1 - use the previous line read from the SFU for decoding the first line at the start of the next band. 0 - ignore the previous line read from the SFU for decoding the first line at the start of the next band (an all 0's line is used instead). 0x20 NextBandEnable 1 0x0 If (NextBandEnable == 1 & Go == 0) then NextBandCurrReadAdr is copied to CurrReadAdr, NextBandLinesRemaining is copied to LinesRemaining, NextBandPrevLineSource is copied to PrevLineSource, Go is set, NextBandEnable is cleared. To start LBD processing NextBandEnable should be set. Work registers (read only for external access) 0x24 CurrReadAdr[21:5] 17 — The current 256-bit aligned read address (256-bit aligned DRAM within the compressed bi-level image (DRAM address) address). Read only register. 0x28 LinesRemaining 15 — Count of number of lines remaining to be decoded. The band has finished when this number reaches 0. Read only register. 0x2C PrevLineSource 1 — 1 - uses the previous line read from the SFU for decoding the first line at the start of the next band. 0 - ignores the previous line read from the SFU for decoding the first line at the start of the next band (an all 0's line is used instead). Read only register. 0x30 CurrWriteAdr 15 — The current dot position for writing to the SFU. Read only register. 0x34 FirstLineOfBand 1 — Indicates whether the current line is considered to be the first line of the band. Read only register. 24.3.3 Starting the LBD Between Bands

The LBD should be started after the SFU. The LBD is programed with a start address for the compressed bi-level data, a decode line length, the source of the previous line and a count of how many lines to decode. The LBD's NextBandEnable bit should then be set (this will set LBD Go). The LBD decodes a single band and then stops, clearing it's Go bit and issuing a pulse on lbd_finishedband. The LBD can then be restarted for the next band, while the HCU continues to process previously decoded bi-level data from the SFU.

There are 4 mechanisms for restarting the LBD between bands:

-   a. lbd_finishedband causes an interrupt to the CPU. The LBD will     have stopped and cleared its Go bit. The CPU reprograms the LBD,     typically the NextBandCurrReadAdr, NextBandLinesRemaining and     NextBandPrevLineSource shadow registers, and sets NextBandEnable to     restart the LBD. -   b. The CPU programs the LBD's NextBandCurrReadAdr,     NextBandLinesRemaining, and NextBandPrevLineSource shadow registers     and sets the NextBandEnable flag before the end of the current band.     At the end of the band the LBD clears Go, NextBandEnable is already     set so the LBD restarts immediately. -   c. The PCU is programmed so that lbd_finishedband triggers the PCU     to execute commands from DRAM to reprogram the LBD's     NextBandCurrReadAdr, NextBandLinesRemaining, and     NextBandPrevLineSource shadow registers and set NextBandEnable to     restart the LBD. The advantage of this scheme is that the CPU could     process band headers in advance and store the band commands in DRAM     ready for execution. -   d. This is a combination of b and c above. The PCU (rather than the     CPU in b) programs the LBD's NextBandCurrReadAdr,     NextBandLinesRemaining, and NextBandPrevLineSource shadow registers     and sets the NextBandEnable flag before the end of the current band.     At the end of the band the LBD clears Go and pulses     lbd_finishedband. NextBandEnable is already set so the LBD restarts     immediately. Simultaneously, lbd_finishedband triggers the PCU to     fetch commands from DRAM. The LBD will have restarted by the time     the PCU has fetched commands from DRAM. The PCU commands program the     LBD's shadow registers and sets NextBandEnable for the next band.     24.3.4 Top-Level Description

A block diagram of the LBD is shown in FIG. 148.

The LBD contains the following sub-blocks:

TABLE 157 Functional sub-blocks in the LBD name Description Registers and PCU interface and configuration registers. Resets Also generates the Go and the Reset signals for the rest of the LBD Stream Decoder Accesses the bi-level description from the DRAM through the DIU interface. It decodes the bit stream into a command with arguments, which it then passes to the command controller. Command Controller Interprets the command from the stream decoder and provide the line fill unit with a limit address and color to fill the SFU Next Line Buffer. It also provides the next edge unit starting address to look for the next edge. Next Edge Unit Scans through the Previous Line Buffer using its current address to find the next edge of a color provided by the command controller. The next edge unit outputs this as the next current address back to the command con- troller and sets a valid bit when this address is at the next edge. Line Fill Unit Fills the SFU Next Line Buffer with a color from its current address up to a limit address. The color and limit are provided by the command controller.

In the following description the LBD decodes data for its current decode line but writes this data into the SFU's next line buffer.

Naming of signals and logical blocks are taken from [22].

The LBD is able to stall mid-line should the SFU be unable to supply a previous line or receive a current line frame due to band processing latency.

All output control signals from the LBD must always be valid after reset. For example, if the LBD is not currently decoding, lbd_sfu_advline (to the SFU) and lbd_finishedband will always be 0.

24.3.5 Registers and Resets Sub-block Description

Since the CDU, LBD and TE all access the page band store, they share two registers that enable sequential memory accesses to the page band stores to be circular in nature. The CDU chapter lists these two registers. The register descriptions for the LBD are listed in Table.

During initialisation of the LBD, the LineLength and the LinesRemaining configuration values are written to the LBD. The ‘Registers and Resets’ sub-block supplies these signals to the other sub-blocks in the LBD. In the case of LinesRemaining, this number is decremented for every line that is completed by the LBD.

If pass through is used during a band the PassThroughEnable register needs to be programmed and PassThroughDotLength programmed with the length of the compressed bits in pass through mode.

PrevLineSource is programmed during the initialisation of a band, if the previous line supplied for the first line is a valid previous line, a 1 is written to PrevLineSource so that the data is used. If a 0 is written the LBD ignores the previous line information supplied and acts as if it is receiving all zeros for the previous line regardless of what the out of the SFU is.

The ‘Registers and Resets’ sub-block also generates the resets used by the rest of the LBD and the Go bit which tells the LBD that it can start requesting data from the DIU and commence decoding of the compressed data stream.

24.3.6 Stream Decoder Sub-block Description

The Stream Decoder reads the compressed bi-level image from the DRAM via the DIU (single accesses of 256-bits) into a double 256-bit FIFO. The barrel shift register uses the 64-bit word from the FIFO to fill up the empty space created by the barrel shift register as it is shifting it's contents. The bit stream is decoded into a command/arguments pair, which in turn is passed to the command controller.

A dataflow block diagram of the stream decoder is shown in FIG. 149.

24.3.6.1 DecodeC—Decode Command

The DecodeC logic encodes the command from bits 6..0 of the bit stream to output one of three commands: SKIP, VERTICAL and RUNLENGTH. It also provides an output to indicate how many bits were consumed, which feeds back to the barrel shift register.

There is a fourth command, PASS_THROUGH, which is not encoded in bits 6..0, instead it is inferred in a special runlength. If the stream decoder detects a short runlength value, i.e. a number less than 31, encoded as a medium runlength this tell the Stream Decoder that once the horizontal command containing this runlength is decoded completely the LBD enters PASS_THROUGH mode. Following the runlength there will be a number of bits that represent un-compressed data. The LBD will stay in PASS_THROUGH mode until all these bits have been decoded successfully, this will occur once a programmed number of bits is reached or the line ends, which ever comes first.

24.3.6.2 DecodeD—Decode Delta

The DecodeD logic decodes the run length from bits 20..3 of the bit stream. If DecodeC is decoding a vertical command, it will cause DecodeD to put constants of −3 through 3 on its output. The output delta is a 15 bit number, which is generally considered to be positive, but since it needs to only address to 13824 dots for an A4 page and 19488 dots for an A3 page (of 32,768), a 2's complement representation of −3, −2, −1 will work correctly for the data pipeline that follows. This unit also outputs how many bits were consumed.

In the case of PASS_THROUGH mode, DecodeD parses the bits that represent the un-compressed data and this is used by the Line Fill Unit to construct the current line frame.

DecodeD parses the bits at one bit per clock cycle and passes the bit in the less significant bit location of delta to the line fill unit.

DecodeD currently requires to know the color of the run length to decode it correctly as black and white runs are encoded differently. The stream decoder keeps track of the next color based on the current color and the current command.

24.3.6.3 State-machine

This state machine continuously fetches consecutive DRAM data whenever there is enough free space in the FIFO, thereby keeping the barrel shift register full so it can continually decode commands for the command controller. Note in FIG. 149 that each read cycle curr_read_addr is compared to end_of_band_store. If the two are equal, curr_read_addr is loaded with start_of_band_store (circular memory addressing). Otherwise curr_read_addr is simply incremented. start_of_band_store and end_of_band_store need to be programed so that the distance between them is a multiple of the 256-bit DRAM word size.

When the state machine decodes a SKIP command, the state machine provides two SKIP instructions to the command controller.

The RUNLENGTH command has two different run lengths. The two run lengths are passed to the command controller as separate RUNLENGTH instructions. In the first instruction fetch, the first run length is passed, and the state machine selects the DecodeD shift value for the barrel shift. In the second instruction fetch from the command controller another RUNLENGTH instruction is generated and the respective shift value is decoded. This is achieved by forcing DecodeC to output a second RUNLENGTH instruction and the respective shift value is decoded.

For PASS_THROUGH mode, the PASS THROUGH command is issued every time the command controller requests a new command. It does this until all the un-compressed bits have been processed.

24.3.7 Command Controller Sub-block Description

The Command Controller interprets the command from the Stream Decoder and provides the line fill unit with a limit address and color to fill the SFU Next Line Buffer. It provides the next edge unit with a starting address to look for the next edge and is responsible for detecting the end of line and generating the eob_cc signal that is passed to the line fill unit.

A dataflow block diagram of the command controller is shown in FIG. 150. Note that data names such as a0 and b1p are taken from [22], and they denote the reference or starting changing element on the coding line and the first changing element on the reference line to the right of a0 and of the opposite color to a0 respectively.

24.3.7.1 State Machine

The following is an explanation of all the states that the state machine utilizes.

i Start

This is the state that the Command Controller enters when a hard or soft reset occurs or when Go has been de-asserted. This state cannot be left until the reset has been removed, Go has been asserted and the NEU (Next Edge Unit), the SD (Stream Decoder) and the SFU are ready.

ii AWAIT_BUFFER

The NEU contains a buffer memory for the data it receives from the SFU. When the command controller enters this state the NEU detects this and starts buffering data, the command controller is able to leave this state when the state machine in the NEU has entered the NEU_RUNNING state. Once this occurs the command controller can proceed to the PARSE state.

iii PAUSE_CC

During the decode of a line it is possible for the FIFO in the stream decoder to get starved of data if the DRAM is not able to supply replacement data fast enough. Additionally the SFU can also stall mid-line due to band processing latency. If either of these cases occurs the LBD needs to pause until the stream decoder gets more of the compressed data stream from the DRAM or the SFU can receive or deliver new frames. All of the remaining states check if sdvalid goes to zero (this denotes a starving of the stream decoder) or if sfu_lbd_rdy goes to zero and that the LBD needs to pause. PAUSE_CC is the state that the command controller enters to achieve this and it does not leave this state until sdvalid and sfu_lbd_rdy are both asserted and the LBD can recommence decompressing.

iv PARSE

Once the command controller enters the PARSE state it uses the information that is supplied by the stream decoder. The first clock cycle of the state sees the sdack signal getting asserted informing the stream decoder that the current register information is being used so that it can fetch the next command.

When in this state the command controller can receive one of four valid commands:

a) Runlength or Horizontal

For this command the value given as delta is an integer that denotes the number of bits of the current color that must be added to the current line.

Should the current line position, a0, be added to the delta and the result be greater than the final position of the current frame being processed by the Line Fill Unit (only 16 bits at a time), it is necessary for the command controller to wait for the Line Fill Unit (LFU) to process up to that point. The command controller changes into the WAIT_FOR_RUNLENGTH state while this occurs.

When the current line position, a1, and the delta together equal or exceed the LINE_LENGTH, which is programmed during initialisation, then this denotes that it is the end of the current line. The command controller signals this to the rest of the LBD and then returns to the STAR_(T) state.

b) Vertical

When this command is received, it tells the command controller that, in the previous line, it needs to find a change from the current color to opposite of the current color, i.e. if the current color is white it looks from the current position in the previous line for the next time where there is a change in color from white to black. It is important to note that if a black to white change occurs first it is ignored.

Once this edge has been detected, the delta will denote which of the vertical commands to use, refer to Table. The delta will denote where the changing element in the current line is relative to the changing element on the previous line, for a Vertical(2) the new changing element position in the current line will correspond to the two bits extra from changing element position in the previous line.

Should the next edge not be detected in the current frame under review in the NEU, then the command controller enters the WAIT_FOR_NE state and waits there until the next edge is found.

c) Skip

A skip follow the same functionality as to Vertical(0) commands but the color in the current line is not changed as it is been filled out. The stream decoder supplies what looks like two separate skip commands that the command controller treats the same a two Vertical(0) commands and has been coded not to change the current color in this case.

d) Pass Through

When in pass through mode the stream decoder supplies one bit per clock cycle that is uses to construct the current frame. Once pass through mode is completed, which is controlled in the stream decoder, the LBD can recommence normal decompression again. The current color after pass through mode is the same color as the last bit in un-compressed data stream. Pass through mode does not need an extra state in the command controller as each pass through command received from the stream decoder can always be processed in one clock cycle.

v WAIT_FOR_RUNLENGTH

As some RUNLENGTH's can carry over more than one 16-bit frame, this means that the Line Fill Unit needs longer than one clock cycle to write out all the bits represented by the RUNLENGTH.

After the first clock cycle the command controller enters into the WAIT_FOR_RUNLENGTH state until all the RUNLENGTH data has been consumed. Once finished and provided it is not the end of the line the command controller will return to the PARSE state.

vi WAIT_FOR_NE

Similar to the RUNLENGTH commands the vertical commands can sometimes not find an edge in the current 16-bit frame. After the first clock cycle the command controller enters the WAIT_FOR_NE state and remains here until the edge is detected. Provided it is not the end of the line the command controller will return to the PARSE state.

vii FINISH_LINE

At the end of a line the command controller needs to hold its data for the SFU before going back to the STAR_(T) state. Command controller remains in the FINISH_LINE state for one clock cycle to achieve this.

24.3.8 Next Edge Unit Sub-block Description

The Next Edge Unit (NEU) is responsible for detecting color changes, or edges, in the previous line based on the current address and color supplied by the Command Controller. The NEU is the interface to the SFU and it buffers the previous line for detecting an edge. For an edge detect operation the Command Controller supplies the current address, this typically was the location of the last edge, but it could also be the end of a run length. With the current address a color is also supplied and using these two values the NEU will search the previous line for the next edge. If an edge is found the NEU returns this location to the Command Controller as the next address in the current line and it sets a valid bit to tell the Command Controller that the edge has been detected. The Line Fill Unit uses this result to construct the current line. The NEU operates on 16-bit words and it is possible that there is no edge in the current 16 bits in the NEU. In this case the NEU will request more words from the SFU and will keep searching for an edge. It will continue doing this until it finds an edge or reaches the end of the previous line, which is based on the LINE_LENGTH. A dataflow block diagram of the Next Edge unit is shown in FIG. 152.

24.3.8.1 NEU Buffer

The algorithm being employed for decompression is based on the whole previous line and is not delineated during the line. However the Next Edge Unit, NEU, can only receive 16 bits at a time from the SFU. This presents a problem for vertical commands if the edge occurs in the successive frame, but refers to a changing element in the current frame.

To accommodate this the NEU works on two frames at the same time, the current frame and the first 3 bits from the successive frame. This allows for the information that is needed from the previous line to construct the current frame of the current line.

In addition to this buffering there is also buffering right after the data is received from the SFU as the SFU output is not registered. The current implementation of the SFU takes two clock cycles from when a request for a current line is received until it is returned and registered. However when NEU requests a new frame it needs it on the next clock cycle to maintain a decoded rate of 2 bits per clock cycle. A more detailed diagram of the buffer in the NEU is shown in FIG. 153.

The output of the buffer are two 16-bit vectors, use_prev_line_a and use_prev_line_b, that are used to detect an edge that is relevant to the current line being put together in the Line Fill Unit.

24.3.8.2 NEU Edge Detect

The NEU Edge Detect block takes the two 16 bit vectors supplied by the buffer and based on the current line position in the current line, a0, and the current color, sd_color, it will detect if there is an edge relevant to the current frame. If the edge is found it supplies the current line position, b1p, to the command controller and the line fill unit. The configuration of the edge detect is shown in FIG. 154.

The two vectors from the buffer, use_prev_line_a and use_prev_line_b, pass into two sub-blocks, transition_wtob and transition_btow. transition_wtob detects if any white to black transitions occur in the 19 bit vector supplied and outputs a 19-bit vector displaying the transitions. transition_wtob is functionally the same as transition_btow, but it detects white to black transitions.

The two 19-bit vectors produced enter into a multiplexer and the output of the multiplexer is controlled by color_neu. color_neu is the current edge transition color that the edge detect is searching for.

The output of the multiplexer is masked against a 19-bit vector, the mask is comprised of three parts concatenated together: decode_b_ext, decode_b and FIRST_FLU_WRITE.

The output of transition_wtob (and it complement transition_btow) are all the transitions in the 16 bit word that is under review. The decode_b is a mask generated from a0. In bit-wise terms all the bits above and including a0 are 1's and all bits below a0 are 0's. When they are gated together it means that all the transitions below a0 are ignored and the first transition after a0 is picked out as the next edge.

The decode_b block decodes the 4 lsb of the current address (a0) into 16-bit mask bits that control which of the data bits are examined. Table 158 shows the truth table for this block.

TABLE 158 Decode_b truth table input output 0000 1111111111111111 0001 1111111111111110 0010 1111111111111100 0011 1111111111111000 0100 1111111111110000 0101 1111111111100000 0110 1111111111000000 0111 1111111110000000 1000 1111111100000000 1001 1111111000000000 1010 1111110000000000 1011 1111100000000000 1100 1111000000000000 1101 1110000000000000 1110 1100000000000000 1111 1000000000000000

For cases when there is a negative vertical command from the stream decoder it is possible that the edge is in the three lower significant bits of the next frame. The decode_b_ext block supplies the mask so that the necessary bits can be used by the NEU to detect an edge if present, Table 159 shows the truth table of this block.

TABLE 159 Decode_b_ext truth table delta output Vertical(−3) 111 Vertical(−2) 111 Vertical(−1) 011 OTHERS 001

FIRST_FLU_WRITE is only used in the first frame of the current line. 2.2.5 a) in [22] refers to “Processing the first picture element”, in which it states that “The first starting picture element, a0, on each coding line is imaginarily set at a position just before the first picture element, and is regarded as a white picture element”. transition_wtob and transition_btow are set up produce this case for every single frame. However it is only used by the NEU if it is not masked out. This occurs when FIRST_FLU_WRITE is ‘1’ which is only asserted at the beginning of a line. 2.2.5 b) in [22] covers the case of “Processing the last picture element”, this case states that “The coding of the coding line continues until the position of the imaginary changing element situated after the last actual element is coded”. This means that no matter what the current color is the NEU needs to always find an edge at the end of a line. This feature is used with negative vertical commands.

The vector, end_frame, is a “one-hot” vector that is asserted during the last frame. It asserts a bit in the end of line position, as determined by LineLength, and this simulates an edge in this location which is ORed with the transition's vector. The output of this, masked_data, is sent into the encodeB_one_hot block

24.3.8.3 Encode_b_one_hot

The encode_b_one_hot block is the first stage of a two stage process that encodes the data to determine the address of the 0 to 1 transition. Table 160 lists the truth table outlining the functionally required by this block.

TABLE 160 Encode_b_one_hot Truth Table Input output XXXXXXXXXXXXXXXXXX1 0000000000000000001 XXXXXXXXXXXXXXXXX10 0000000000000000010 XXXXXXXXXXXXXXXX100 0000000000000000100 XXXXXXXXXXXXXXX1000 0000000000000001000 XXXXXXXXXXXXXX10000 0000000000000010000 XXXXXXXXXXXXX100000 0000000000000100000 XXXXXXXXXXXX1000000 0000000000001000000 XXXXXXXXXXX10000000 0000000000010000000 XXXXXXXXXX100000000 0000000000100000000 XXXXXXXXX1000000000 0000000001000000000 XXXXXXXX10000000000 0000000010000000000 XXXXXXX100000000000 0000000100000000000 XXXXXX1000000000000 0000001000000000000 XXXXX10000000000000 0000010000000000000 XXXX100000000000000 0000100000000000000 XXX1000000000000000 0001000000000000000 XX10000000000000000 0010000000000000000 X100000000000000000 0100000000000000000 1000000000000000000 1000000000000000000 0000000000000000000 0000000000000000000

The output of encode_b_one-hot is a “one-hot” vector that will denote where that edge transition is located. In cases of multiple edges, only the first one will be picked.

24.3.8.4 Encode_b_(—)4 Bit

Encode_b_(—)4 bit is the second stage of the two stage process that encodes the data to determine the address of the 0 to 1 transition.

Encode_b_(—)4 bit receives the “one-hot” vector from encode_b_one_hot and determines the bit location that is asserted. If there is none present this means that there was no edge present in this frame. If there is a bit asserted the bit location in the vector is converted to a number, for example if bit 0 is asserted then the number is one, if bit one is asserted then the number is one, etc. The delta supplied to the NEU determines what vertical command is being processed. The formula that is implemented to return b1p to the command controller is: for V(n) bip=x+n modulus16

-   -   where x is the number that was extracted from the “one-hot”         vector and n is the vertical command.         24.3.8.5 State Machine

The following is an explanation of all the states that the NEU state machine utilizes.

i NEU_START

This is the state that NEU enters when a hard or soft reset occurs or when Go has been de-asserted. This state can not left until the reset has been removed, Go has been asserted and it detects that the command controller has entered it's AWAIT_BUFF state. When this occurs the NEU enters the NEU_FILL_BUFF state.

ii NEU_FILL_BUFF

Before any compressed data can be decoded the NEU needs to fill up its buffer with new data from the SFU. The rest of the LBD waits while the NEU retrieves the first four frames from the previous line. Once completed it enters the NEU_HOLD state.

iii NEU_HOLD

The NEU waits in this state for one clock cycle while data requested from the SFU on the last access returns.

iv NEU_RUNNING

NEU_RUNNING controls the requesting of data from the SFU for the remainder of the line by pulsing lbd_sfu_pladvword when the LBD needs a new frame from the SFU. When the NEU has received all the word it needs for the current line, as denoted by the LineLength, the NEU enters the NEU_EMPTY state.

v NEU_EMPTY

NEU waits in this state while the rest of the LBD finishes outputting the completed line to the SFU. The NEU leaves this state when Go gets deasserted. This occurs when the end_of_line signal is detected from the LBD.

24.3.9 Line Fill Unit Sub-block Description

The Line Fill Unit, LFU, is responsible for filling the next line buffer in the SFU. The SFU receives the data in blocks of sixteen bits. The LFU uses the color and a0 provided by the Command Controller and when it has put together a complete 16-bit frame, it is written out to the SFU. The LBD signals to the SFU that the data is valid by strobing the lbd_sfu_wdatavalid signal.

When the LFU is at the end of the line for the current line data it strobes lbd_sfu_advline to indicate to the SFU that the end of the line has occurred.

A dataflow block diagram of the line fill unit is shown in FIG. 154.

The dataflow above has the following blocks:

24.3.9.1 State Machine

The following is an explanation of all the states that the LFU state machine utilizes.

i LFU_START

This is the state that the LFU enters when a hard or soft reset occurs or when Go has been de-asserted. This state can not left until the reset has been removed, Go has been asserted and it detects that a0 is no longer zero, this only occurs once the command controller start processing data from the Next Edge Unit, NEU.

ii LFU_NEW REG

LFU_NEW_REG is only entered at the beginning of a new frame. It can remain in this state on subsequent cycles if a whole frame is completed in one clock cycle. If the frame is completed the LFU will output the data to the SFU with the write enable signal. However if a frame is not completed in one clock cycle the state machine will change to the LFU_COMPLETE_REG state to complete the remainder of the frame. LFU_NEW_REG handles all the lbd_sfu_wdata writes and asserts lbd_sfu_wdatavalid as necessary.

iii LFU_COMPLETE_REG

LFU_COMPLETE_REG fills out all the remaining parts of the frame that were not completed in the first clock cycle. The command controller supplies the a0 value and the color and the state machine uses these to derive the limit and color_sel_(—)16 bit_If which the line_fill_data block needs to construct a frame. Limit is the four lower significant bits of a0 and color_sel_(—)16 bit_If is a 16-bit wide mask of sd_dolor. The state machine also maintains a check on the upper eleven bits of a0. If these increment from one clock cycle to the next that means that a frame is completed and the data can be written to the SFU. In the case of the LineLength being reached the Line Fill Unit fills out the remaining part of the frame with the color of the last bit in the line that was decoded.

24.3.9.2 line_fill_data

line_fill_data takes the limit value and the color_sel_(—)16 bit_If values and constructs the current frame that the command controller and the next edge unit are decoding. The following pseudo code illustrate the logic followed by the line_fill_data. work_sfu_wdata is exported by the LBD to the SFU as lbd_sfu_wdata.

if (lfu_state == LFU_START) OR (lfu_state == LFU_NEW_REG) then work_sfu_wdata = color_sel_16bit_lf else work_sfu_wdata[(15 − limit) downto limit] = color_sel_16bit_lf[(15 − limit) downto limit] 25 Spot FIFO Unit (SFU) 25.1 Overview

The Spot FIFO Unit (SFU) provides the means by which data is transferred between the LBD and the HCU. By abstracting the buffering mechanism and controls from both units, the interface is clean between the data user and the data generator. The amount of buffering can also be increased or decreased without affecting either the LBD or HCU. Scaling of data is performed in the horizontal and vertical directions by the SFU so that the output to the HCU matches the printer resolution. Non-integer scaling is supported in both the horizontal and vertical directions. Typically, the scale factor will be the same in both directions but may be programmed to be different.

25.2 Main Features of the SFU

The SFU replaces the Spot Line Buffer Interface (SLBI) in PEC1. The spot line store is now located in DRAM.

The SFU outputs the previous line to the LBD, stores the next line produced by the LBD and outputs the HCU read line. Each interface to DRAM is via a feeder FIFO. The LBD interfaces to the SFU with a data width of 16 bits. The SFU interfaces to the HCU with a data width of 1 bit. Since the DRAM word width is 256-bits but the LBD line length is a multiple of 16 bits, a capability to flush the last multiples of 16-bits at the end of a line into a 256-bit DRAM word size is required. Therefore, SFU reads of DRAM words at the end of a line, which do not fill the DRAM word, will already be padded.

A signal sfu_lbd_rdy to the LBD indicates that the SFU is available for writing and reading. For the first LBD line after SFU Go has been asserted, previous line data is not supplied until after the first lbd_sfu_advline strobe from the LBD (zero data is supplied instead), and sfu_lbd_rdy to the LBD indicates that the SFU is available for writing. lbd_sfu_advline tells the SFU to advance to the next line. lbd_sfu_pladvword tells the SFU to supply the next 16-bits of previous line data. Until the number of lbd_sfu_pladvword strobes received is equivalent to the LBD line length, sfu_lbd_rdy indicates that the SFU is available for both reading and writing. Thereafter it indicates the SFU is available for writing. The LBD should not generate lbd_sfu_pladvword or lbd_sfu_advline strobes until sfu_lbd_rdy is asserted.

A signal sfu_hcu_avail indicates that the SFU has data to supply to the HCU. Another signal hcu_sfu_advdot, from the HCU, tells the SFU to supply the next dot. The HCU should not generate the hcu_sfu_advdot signal until sfu_hcu_avail is true. The HCU can therefore stall waiting for the sfu_hcu_avail signal.

X and Y non-integer scaling of the bi-level dot data is performed in the SFU.

At 1600 dpi the SFU requires 1 dot per cycle for all DRAM channels, 3 dots per cycle in total (read+read+write). Therefore the SFU requires two 256 bit read DRAM access per 256 cycles, 1 write access every 256 cycles. A single DIU read interface will be shared for reading the current and previous lines from DRAM.

25.3 Bi-Level Dram Memory Buffer Between LBD, SFU and HCU

FIG. 158 shows a bi-level buffer store in DRAM. FIG. 158( a) shows the LBD previous line address reading after the HCU read line address in DRAM. FIG. 158( b) shows the LBD previous line address reading before the HCU read line address in DRAM.

Although the LBD and HCU read and write complete lines of data, the bi-level DRAM buffer is not line based. The buffering between the LBD, SFU and HCU is a FIFO of programmable size. The only line based concept is that the line the HCU is currently reading cannot be over-written because it may need to be re-read for scaling purposes.

The SFU interfaces to DRAM via three FIFOs:

-   a. The HCUReadLineFIFO which supplies dot data to the HCU. -   b. The LBDNextLineFIFO which writes decompressed bi-level data from     the LBD. -   c. The LBDPrevLineFIFO which reads previous decompressed bi-level     data for the LBD.

There are four address pointers used to manage the bi-level DRAM buffer:

-   a. hcu_readline_rd_adr[21:5] is the read address in DRAM for the     HCUReadLineFIFO. -   b. hcu_startreadline_adr[21:5] is the start address in DRAM for the     current line being read by the HCUReadLineFIFO. -   c. lbd_nextline_wr_adr[21:5] is the write address in DRAM for the     LBDNextLineFIFO. -   d. lbd_prevline_rd_adr[21:5] is the read address in DRAM for the     LBDPrevLineFIFO.

The address pointers must obey certain rules which indicate whether they are valid:

-   a. hcu_readline_rd_adr is only valid if it is reading earlier in the     line than lbd_nextline_wr_adr is writing i.e. the fifo is not empty -   b. The SFU (lbd_nextline_wr_adr) cannot overwrite the current line     that the HCU is reading from (hcu_startreadline_adr) i.e. the fifo     is not full, when compared with the HCU read line pointer -   c. The LBDNextLineFIFO (lbd_nextline_wr_adr) must be writing earlier     in the line than LBDPrevLineFIFO (lbd_prevline_rd_adr) is reading     and must not overwrite the current line that the HCU is reading from     i.e. the fifo is not full when compared to the PrevLineFifo read     pointer -   d. The LBDPrevLineFIFO (lbd_prevline_rd_adr) can read right up to     the address that LBDNextLineFIFO (lbd_nextline_wr_adr) is writing     i.e the fifo is not empty. -   e. At startup i.e. when sfu_go is asserted, the pointers are reset     to start_sfu_adr[21:5]. -   f. The address pointers can wrap around the SFU bi-level store area     in DRAM.

As a guideline, the typical FIFO size should be a minimum of 2 lines stored in DRAM, nominally 3 lines, up to a programmable number of lines. A larger buffer allows lines to be decompressed in advance. This can be useful for absorbing local complexities in compressed bi-level images.

25.4 DRAM Access Requirements

The SFU has 1 read interface to the DIU and 1 write interface. The read interface is shared between the previous and current line read FIFOs.

The spot line store requires 5.1 Kbytes of DRAM to store 3 A4 lines. The SFU will read and write the spot line store in single 256-bit DRAM accesses. The SFU will need 256-bit double buffers for each of its previous, current and next line interfaces.

The SFU's DIU bandwidth requirements are summarized in Table 161.

TABLE 161 DRAM bandwidth requirements Maximum number of Peak Bandwidth requir- Average Direc- cycles between each ed to be supported by Bandwidth tion 256-bit DRAM access DIU (bits/cycle) (bits/cycle) Read 1281 2 2 Write 2562 1 1 1: Two separate reads of 1 bit/cycle. 2: Write at 1 bit/cycle. 25.5 Scaling

Scaling of bi-level data is performed in both the horizontal and vertical directions by the SFU so that the output to the HCU matches the printer resolution. The SFU supports non-integer scaling with the scale factor represented by a numerator and a denominator. Only scaling up of the bi-level data is allowed, i.e. the numerator should be greater than or equal to the denominator. Scaling is implemented using a counter as described in the pseudocode below. An advance pulse is generated to move to the next dot (x-scaling) or line (y-scaling).

if (count + denominator >= numerator) then count = (count + denominator) − numerator advance = 1 else count = count + denominator advance = 0

X scaling controls whether the SFU supplies the next dot or a copy of the current dot when the HCU asserts hcu_sfu_advdot. The SFU counts the number of hcu_sfu_advdot signals from the HCU. When the SFU has supplied an entire HCU line of data, the SFU will either re-read the current line from DRAM or advance to the next line of HCU read data depending on the programmed Y scale factor.

An example of scaling for numerator=7 and denominator=3 is given in Table 162. The signal advance if asserted causes the next input dot to be output on the next cycle, otherwise the same input dot is output

TABLE 162 Non-integer scaling example for scaleNum = 7, scaleDenom = 3 count advance dot 0 0 1 3 0 1 6 1 1 2 0 2 5 1 2 1 0 3 4 1 3 0 0 4 3 0 4 6 1 4 2 0 5 25.6 Lead-In and Lead-Out Clipping

To account for the case where there may be two SoPEC devices, each generating its own portion of a dot-line, the first dot in a line may not be replicated the total scale-factor number of times by an individual SoPEC. The dot will ultimately be scaled-up correctly with both devices doing part of the scaling, one on its lead-out and the other on its lead in. Scaled up dots on the lead-out, i.e. which go beyond the HCU linelength, will be ignored. Scaling on the lead-in, i.e. of the first valid dot in the line, is controlled by setting the XstartCount register.

At the start of each line count in the pseudo-code above is set to XstartCount. If there is no lead-in, XstartCount is set to 0 i.e. the first value of count in Table. If there is lead-in then XstartCount needs to be set to the appropriate value of count in the sequence above.

25.7 Interfaces Between LDB, SFU and HCU

25.7;1 LDB-SFU Interfaces

The LBD has two interfaces to the SFU. The LBD writes the next line to the SFU and reads the previous line from the SFU.

25.7.1.1 LBDNextLineFIFO Interface

The LBDNextLineFIFO interface from the LBD to the SFU comprises the following signals:

-   -   lbd_sfu_wdata, 16-bit write data.     -   lbd_sfu_wdatavalID, write data valid.     -   lbd_sfu_advline, signal indicating LDB has advanced to the next         line.

The LBD should not write to the SFU until sfu_lbd_rdy is true. The LBD can therefore stall waiting for the sfu_lbd_rdy signal.

25.7.1.2 LBDPrevLineFIFO Interface

The LBDPrevLineFIFO interface from the SFU to the LBD comprises the following signals:

-   -   sfu_lbd_pldata, 16-bit data.

The previous line read buffer interface from the LBD to the SDU comprises the following signals:

-   -   lbd_sfu_pladvword, signal indicating to the SFU to supply the         next 16-bit word.     -   lbd_sfu_advline, signal indicating LDB has advanced to the next         line.

Previous line data is not supplied until after the first lbd_sfu_advline strobe from the LBD (zero data is supplied instead). The LBD should not assert lbd_sfu_pladvword unless sfu_lbd_rdy is asserted.

25.7.1.3 Common Control Signals

sfu_lbd_rdy indicates to the LBD that the SFU is available for writing. After the first lbd_sfu_advline and before the number of lbd_sfu_pladvword strobes received is equivalent to the LBD line length, sfu_lbd_rdy indicates that the SFU is available for both reading and writing.

Thereafter it indicates the SFU is available for writing.

The LBD should not generate lbd_sfu_pladvword or lbd_sfu_advline strobes until sfu_lbd_rdy is asserted.

25.7.2 SFU-HCU Current Line FIFO Interface

The interface from the SFU to the HCU comprises the following signals:

-   -   sfu_hcu_sdata, 1-bit data.     -   sfu_hcu_avail, data valid signal indicating that there is data         available in the SFU HCUReadLineFIFO.

The interface from HCU to SFU comprises the following signals:

-   -   hcu_sfu_advdot, indicating to the SFU to supply the next dot.

The HCU should not generate the hcu_sfu_advdot signal until sfu_hcu_avail is true. The HCU can therefore stall waiting for the sfu_hcu_avail signal.

25.8 Implementation

25.8.1 Definitions of IO

TABLE 163 SFU Port List Port Name Pins I/O Description Clocks and Resets Pclk 1 In SoPEC Functional clock. prst_n 1 In Global reset signal. DIU Read Interface signals sfu_diu_rreq 1 Out SFU requests DRAM read. A read request must be accompanied by a valid read address. sfu_diu_radr[21:5] 17 Out Read address to DIU 17 bits wide (256-bit aligned word). diu_sfu_rack 1 In Acknowledge from DIU that read request has been accepted and new read address can be placed on sfu_diu_radr. diu_data[63:0] 64 In Data from DIU to SoPEC Units. First 64-bits are bits 63:0 of 256 bit word. Second 64-bits are bits 127:64 of 256 bit word. Third 64-bits are bits 191:128 of 256 bit word. Fourth 64-bits are bits 255:192 of 256 bit word. diu_sfu_rvalid 1 In Signal from DIU telling SoPEC Unit that valid read data is on the diu_data bus. DIU Write Interface signals sfu_diu_wreq 1 Out SFU requests DRAM write. A write request must be accompanied by a valid write address together with valid write data and a write valid. sfu_diu_wadr[21:5] 17 Out Write address to DIU 17 bits wide (256-bit aligned word). diu_sfu_wack 1 In Acknowledge from DIU that write request has been accepted and new write address can be placed on sfu_diu_wadr. sfu_diu_data[63:0] 64 Out Data from SFU to DIU. First 64-bits are bits 63:0 of 256 bit word. Second 64-bits are bits 127:64 of 256 bit word. Third 64-bits are bits 191:128 of 256 bit word. Fourth 64-bits are bits 255:192 of 256 bit word. sfu_diu_wvalid 1 Out Signal from PEP Unit indicating that data on sfu_diu_data is valid. PCU Interface data and control signals pcu_adr[5:2] 4 In PCU address bus. Only 4 bits are required to decode the address space for this block pcu_dataout[31:0] 32 In Shared write data bus from the PCU sfu_pcu_datain[31:0] 32 Out Read data bus from the SFU to the PCU pcu_rwn 1 In Common read/not-write signal from the PCU pcu_sfu_sel 1 In Block select from the PCU. When pcu_sfu_sel is high both pcu_adr and pcu_dataout are valid sfu_pcu_rdy 1 Out Ready signal to the PCU. When sfu_pcu_rdy is high it indicates the last cycle of the access. For a write cycle this means pcu_dataout has been registered by the block and for a read cycle this means the data on sfu_pcu_datain is valid. LBD Interface Data and Control Signals sfu_lbd_rdy 1 Out Signal indication that SFU has previous line data available and is ready to be written to. lbd_sfu_advline 1 In Line advance signal for both next and previous lines. lbd_sfu_pladvword 1 In Advance word signal for previous line buffer. sfu_lbd_pldata[15:0] 16 Out Data from the previous line buffer. lbd_sfu_wdata[15:0] 16 In Write data for next line buffer. lbd_sfu_wdatavalid 1 In Write data valid signal for next line buffer data. HCU Interface Data and Control Signals hcu_sfu_advdot 1 In Signal indicating to the SFU that the HCU is ready to accept the next dot of data from SFU. sfu_hcu_sdata 1 Out Bi-level dot data. sfu_hcu_avail 1 Out Signal indicating valid bi-level dot data on sfu_hcu_sdata. 25.8.2 Configuration Registers

TABLE 164 SFU Configuration Registers Address # value on (SFU_base+) register name bits reset description Control registers 0x00 Reset 1 0x1 A write to this register causes a reset of the SFU. This register can be read to indicate the reset state: 0 - reset in progress 1 - reset not in progress 0x04 Go 1 0x0 Writing 1 to this register starts the SFU. Writing 0 to this register halts the SFU. When Go is deasserted the state- machines go to their idle states but all counters and configuration registers keep their values. When Go is asserted all counters are reset, but configuration registers keep their values (i.e. they don't get reset). The SFU must be started before the LBD is started. This register can be read to determine if the SFU is running (1 - running, 0 - stopped). Setup registers (constant for during processing the page) 0x08 HCUNumDots 16 0x0000 Width of HCU line (in dots). 0x0C HCUDRAMWords 8 0x00 Number of 256-bit DRAM words in a HCU line − 1. 0x10 LBDDRAMWords 8 0x00 Number of 256-bit words in a LBD line − 1. (LBD line length must be at least 128 bits). 0x14 StartSfuAdr[21:5] 17 0x0000 0 First SFU location in memory. (256-bit aligned DRAM address) 0x18 EndSfuAdr[21:5] 17 0x0000 0 Last SFU location in memory. (256-bit aligned DRAM address) 0x1C XstartCount 8 0x00 Value to be loaded at the start of every line into the counter used for scaling in the X direction. Used to control the scaling of the first dot in a line. This value will typically equal zero, except in the case where a number of dots are clipped on the lead in to a line. XstartCount must be programmed to be less than the XscaleNum value. 0x20 XscaleNum 8 0x01 Numerator of spot data scale factor in X direction. 0x24 XscaleDenom 8 0x01 Denominator of spot data scale factor in X direction. 0x28 YscaleNum 8 0x01 Numerator of spot data scale factor in Y direction. 0x2C YscaleDenom 8 0x01 Denominator of spot data scale factor in Y direction. Work registers (PCU has read-only access) 0x30 HCUReadLineAdr[21:5] 17 — Current address pointer in DRAM to (256-bit aligned HCU read data. Read only register. DRAM address) 0x34 HCUStartReadLineAdr[21:5] 17 — Start address in DRAM of line being (256-bit aligned read by HCU buffer in DRAM. Read DRAM address) only register. 0x38 LBDNextLineAdr[21:5] 17 — Current address pointer in DRAM to (256-bit aligned LBD write data. Read only register DRAM address) 0x3C LBDPrevLineAdr[21:5] 17 — Current address pointer in DRAM to (256-bit aligned LBD read data. Read only register DRAM address) 25.8.3 SFU Sub-block Partition

The SFU contains a number of sub-blocks:

Name description PCU Interface PCU interface, configuration and status registers. Also generates the Go and the Reset signals for the rest of the SFU LBD Previous Contains FIFO which is read by the LBD Line FIFO previous line interface. LBD Next Contains FIFO which is written by the LBD Line FIFO next line interface. HCU Read Contains FIFO which is read by the HCU Line FIFO interface. DIU Interface Contains DIU read interface and DIU write and Address interface. Manages the address pointers for Generator the bi-level DRAM buffer. Contains X and Y scaling logic.

The various FIFO sub-blocks have no knowledge of where in DRAM their read or write data is stored. In this sense the FIFO sub-blocks are completely de-coupled from the bi-level DRAM buffer. All DRAM address management_is centralised in the DIU Interface and Address Generation sub-block. DRAM access is pre-emptive i.e. after a FIFO unit has made an access then as soon as the FIFO has space to read or data to write a DIU access will be requested immediately. This ensures there are no unnecessary stalls introduced e.g. at the end of an LBD or HCU line.

There now follows a description of the SFU sub-blocks.

25.8.4 PCU Interface Sub-block

The PCU interface sub-block provides for the CPU to access SFU specific registers by reading or writing to the SFU address space.

25.8.5 LBDPrevLineFIFO Sub-block

TABLE 165 LBDPrevLineFIFO Additional IO Definitions Port Name Pins I/O Description Internal Output plf_rdy 1 Out Signal indicating LBDPrevLineFIFO is ready to be read from. Until the first lbd_sfu_advline for a band has been received and after the number of reads from DRAM for a line is received is equal to LBDDRAMWords, plf_rdy is always asserted. During the second and subsequent lines plf_rdy is deasserted whenever the LBDPrevLineFIFO has one word left in the FIFO.. DIU and Address Generation sub-block Signals plf_diurreq 1 Out Signal indicating the LBDPrevLineFIFO has 256-bits of data free. plf_diurack 1 In Acknowledge that read request has been accepted and plf_diurreq should be de-asserted. plf_diurdata 1 In Data from the DIU to LBDPrevLineFIFO. First 64-bits are bits 63:0 of 256 bit word. Second 64-bits are bits 127:64 of 256 bit word. Third 64-bits are bits 191:128 of 256 bit word. Fourth 64-bits is are 255:192 of 256 bit word. plf_diurrvalid 1 In Signal indicating data on plf_diurdata is valid. plf_diuidle 1 Out Signal indicating DIU state-machine is in the IDLE state. 25.8.5.1 General Description

The LBDPrevLineFIFO sub-block comprises a double 256-bit buffer between the LBD and the DIU Interface and Address Generator sub-block. The FIFO is implemented as 8 times 64-bit words. The FIFO is written by the DIU Interface and Address Generator sub-block and read by the LBD.

Whenever 4 locations in the FIFO are free the FIFO will request 256-bits of data from the DIU Interface and Address Generation sub-block by asserting plf_diurreq. A signal plf_diurack indicates that the request has been accepted and plf_diurreq should be de-asserted.

The data is written to the FIFO as 64-bits on plf_diurdata[63:0] over 4 clock cycles. The signal plf_diurvalid indicates that the data returned on plf_diurdata[63:0] is valid. plf_diurvalid is used to generate the FIFO write enable, write_en, and to increment the FIFO write address, write_adr[2:0]. If the LBDPrevLineFIFO still has 256-bits free then plf_diurreq should be asserted again.

The DIU Interface and Address Generation sub-block handles all address pointer management and DIU interfacing and decides whether to acknowledge a request for data from the FIFO.

The state diagram of the LBDPrevLineFIFO DIU Interface is shown in FIG. 163. If sfu_go is deasserted then the state-machine returns to its idle state.

The LBD reads 16-bit wide data from the LBDPrevLineFIFO on sfu_lbd_pldata[15:0]. lbd_sfu_pladvword from the LBD tells the LBDPrevLineFIFO to supply the next 16-bit word. The FIFO control logic generates a signal word_select which selects the next 16-bits of the 64-bit FIFO word to output on sfu_lbd_pldata[15:0]. When the entire current 64-bit FIFO word has been read by the LBD lbd_sfu_pladvword will cause the next word to be popped from the FIFO. Previous line data is not supplied until after the first lbd_sfu_advline strobe from the LBD after sfu_go is asserted (zero data is supplied instead). Until the first lbd_sfu_advline strobe after sfu_go lbd_sfu_pladvword strobes are ignored.

The LBDPrevLineFIFO control logic uses a counter, pl_count[7:0], to counts the number of DRAM read accesses for the line. When the pl_count counter is equal to the LBDDRAMWords, a complete line of data has been read by the LBD the plf_rdy is set high, and the counter is reset. It remains high until the next lbd_sfu_advline strobe from the LBD. On receipt of the lbd_sfu_advline strobe the remaining data in the 256-bit word in the FIFO is ignored, and the FIFO read_adr is rounded up if required.

The LBDPrevLineFIFO generates a signal plf_rdy to indicate that it has data available. Until the first lbd_sfu_advline for a band has been received and after the number of DRAM reads for a line is equal to LBDDRAMWords, plf_rdy is always asserted. During the second and subsequent lines plf_rdy is deasserted whenever the LBDPrevLineFIFO has one word left.

The last 256-bit word for a line read from DRAM can contain extra padding which should not be output to the LBD. This is because the number of 16-bit words per line may not fit exactly into a 256-bit DRAM word. When the count of the number of DRAM reads for a line is equal to lbd_dram_words the LBDPrevLineFIFO must adjust the FIFO write address to point to the next 256-bit word boundary in the FIFO for the next line of data. At the end of a line the read address must round up the nearest 256-bit word boundary and ignore the remaining 16-bit words. This can be achieved by considering the FIFO read address, read adr[2:0], will require 3 bits to address 8 locations of 64-bits. The next 256-bit aligned address is calculated by inverting the MSB of the read_adr and setting all other bits to 0.

if (read_adr[1:0] /= b00 AND lbd_sfu_advline == 1)then read_adr[1:0] = b00 read_adr[2] = ~read_adr[2] 25.8.6 LBDNextLineFIFO Sub-block

TABLE 166 LBDNextLineFIFO Additional IO Definition Port Name Pins I/O Description LBDNextLineFIFO Interface Signals nlf_rdy 1 Out Signal indicating LBDNextLineFIFO is ready to be written to i.e. there is space in the FIFO. DIU and Address Generation sub-block Signals nlf_diuwreq 1 Out Signal indicating the LBDNextLineFIFO has 256-bits of data for writing to the DIU. nlf_diuwack 1 In Acknowledge from DIU that write request has been accepted and write data can be output on nlf_diuwdata together with nlf_diuwvalid. nlf_diuwdata 1 Out Data from LBDNextLineFIFO to DIU Interface. First 64-bits is bits 63:0 of 256 bit word Second 64-bits is bits 127:64 of 256 bit word Third 64-bits is bits 191:128 of 256 bit word Fourth 64-bits is bits 255:192 of 256 bit word nlf_diuwvalid 1 In Signal indicating that data on wlf_diuwdata is valid. 25.8.6.1 General Description

The LBDNextLineFIFO sub-block comprises a double 256-bit buffer between the LBD and the DIU Interface and Address Generator sub-block. The FIFO is implemented as 8 times 64-bit words. The FIFO is written by the LBD and read by the DIU Interface and Address Generator. Whenever 4 locations in the FIFO are full the FIFO will request 256-bits of data to be written to the DIU Interface and Address Generator by asserting nlf_diuwreq. A signal nlf_diuwack indicates that the request has been accepted and nlf_diuwreq should be de-asserted. On receipt of nlf_diuwack, the data is sent to the DIU Interface as 64-bits on nlf_diuwdata[63:0] over 4 clock cycles. The signal nlf_diuwvalid indicates that the data on nlf_diuwdata[63:0] is valid. nlf_diuwvalid should be asserted with the smallest latency after nlf_diuwack. If the LBDNextLineFIFO still has 256-bits more to transfer then nlf_diuwreq should be asserted again. The state diagram of the LBDNextLineFIFO DIU Interface is shown in FIG. 166. If sfu_go is deasserted then the state-machine returns to its Idle state.

The signal nlf_rdy indicates that the LBDNextLineFIFO has space for writing by the LBD. The LBD writes 16-bit wide data supplied on lbd_sfu_wdata[15:0]. lbd_sfu_wvalid indicates that the data is valid.

The LBDNextLineFIFO control logic counts the number of lbd_sfu_wvalid signals and is used to correctly address into the next line FIFO. The lbd_sfu_wvalid counter is rounded up to the nearest 256-bit word when a lbd_sfu_advline strobe is received from the LBD. Any data remaining in the FIFO is flushed to DRAM with padding being added to fill a complete 256-bit word.

25.8.7 sfu_lbd_rdy Generation

The signal sfu_lbd_rdy is generated by ANDing plf_rdy from the LBDPrevLineFIFO and nlf_rdy from the LBDNextLineFIFO.

sfu_lbd_rdy indicates to the LBD that the SFU is available for writing i.e. there is space available in the LBDNextLineFIFO. After the first ibd_sfu_advline and before the number of lbd_sfu_pladvword strobes received is equivalent to the line length, sfu_lbd_rdy indicates that the SFU is available for both reading, i.e. there is data in the LBDPrevLineFIFO, and writing.

Thereafter it indicates the SFU is available for writing.

25.8.8 LBD-SFU Interfaces Timing Waveform Description

In FIG. 167 and FIG. 168, shows the timing of the data valid and ready signals between the SFU and LBD. A diagram and pseudocode is given for both read and write interfaces between the SFU and LBD.

25.8.8.1 LBD-SFU Write Interface Timing

The main points to note from FIG. 167 are:

-   -   In clock cycle 1 sfu_lbd_rdy detects that it has only space to         receive 2 more 16 bit words from the LBD after the current clock         cycle.     -   The data on lbd_sfu_wdata is valid and this is indicated by         lbd_sfu_wdatavalid being asserted.     -   In clock cycle 2 sfu_lbd_rdy is deasserted however the LBD can         not react to this signal until clock cycle 3. So in clock cycle         3 there is also valid data from the LBD which consumes the last         available location available in the FIFO in the SFU (FIFO free         level is zero).     -   In clock cycle 4 and 5 the FIFO is read and 2 words become free         in the FIFO.     -   In cycle 4 the SFU determines that the FIFO has more room and         asserts the ready signal on the next cycle.     -   The LBD has entered a pause mode and waits for sfu_lbd_rdy to be         asserted again, in cycle 5 the LBD sees the asserted ready         signal and responds by writing one unit into the FIFO, in cycle         6.     -   The SFU detects it has 2 spaces left in the FIFO and the current         cycle is an active write (same as in cycle 1), and deasserts the         ready on the next cycle.     -   In cycle 7 the LBD did not have data to write into the FIFO, and         so the FIFO remains with one space left     -   The SFU toggles the ready signal every second cycle, this allows         the LBD to write one unit at a time to the FIFO.     -   In cycle 9 the LBD responds to the single ready pulse by writing         into the FIFO and consuming the last remaining unit free.

The write interface pseudocode for generating the ready is.

// ready generation pseudocode if (fifo_free_level > 2) then nlf_rdy = 1 elsif (fifo_free_level == 2) then if (lbd_sfu_wdatavalid == 1)then nlf_rdy = 0 else nlf_rdy = 1 elsif (fifo_free_level == 1) then if (lbd_sfu_wdatavalid == 1)then nlf_rdy = 0 else nlf_rdy = NOT(sfu_lbd_rdy) else nlf_rdy = 0 sfu_lbd_rdy = (nlf_rdy AND plf_rdy) 25.8.8.2 SFU-LBD Read Interface

The read interface is similar to the write interface except that read data (sfu_lbd_pldata) takes an extra cycle to respond to the data advance signal (lbd_sfu_pladvword signal).

It is not possible to read the FIFO totally empty during the processing of a line, one word must always remain in the FIFO. At the end of a line the fifo can be read to totally empty. This functionality is controlled by the SFU with the generation of the plf_signal.

There is an apparent corner case on the read side which should be highlighted. On examination this turns out to not be an issue.

Scenario 1:

-   -   sfu_lbd_rdy will go low when there is still is still 2 pieces of         data in the FIFO. If there is a lbd_sfu_pladvword pulse in the         next cycle the data will appear on sfu_lbd_pldata[15:0].         Scenario 2:     -   sfu_lbd_rdy will go low when there is still 2 pieces of data in         the FIFO. If there is no lbd_sfu_pladvword pulse in the next         cycle and it is not the end of the page then the SFU will read         the data for the next line from DRAM and the read FIFO will fill         more, sfu_lbd_rdy will assert again, and so the data will appear         on sfu_lbd_pldata[15:0]. If it happens that the next line of         data is not available yet the sfu_lbd_pldata bus will go invalid         until the next lines data is available. The LBD does not sample         the sfu_lbd_pldata bus at this time (i.e. after the end of a         line) and it is safe to have invalid data on the bus.         Scenario 3:     -   sfu_lbd_rdy will go low when there is still 2 pieces of data in         the FIFO. If there is no lbd_sfu_pladvword pulse in the next         cycle and it is the end of the page then the SFU will do no more         reads from DRAM, sfu_lbd_rdy will remain de-asserted, and the         data will not be read out from the FIFO. However last line of         data on the page is not needed for decoding in the LBD and will         not be read by the LBD. So scenario 3 will never apply.

The pseudocode for the read FIFO ready generation

// ready generation pseudocode if (pl_count == lbd_dram_words) then plf_rdy = 1 elsif (fifo_fill_level > 3)then plf_rdy = 1 elsif (fifo_fill_level == 3) then if (lbd_sfu_pladvword == 1)then plf_rdy = 0 else plf_rdy = 1 elsif (fifo_fill_level == 2) then if (lbd_sfu_pladvword == 1)then plf_rdy = 0 else plf_rdy = NOT(sfu_lbd_rdy) else plf_rdy = 0 sfu_lbd_rdy = (plf_rdy AND nlf_rdy) 25.8.9 HCUReadLineFIFO Sub-Block

TABLE 167 HCUReadLineFIFO Additional IO Definition Port Name Pins I/O Description DIU and Address Generation sub-block Signals hrf_xadvance 1 In Signal from horizontal scaling unit 1 - supply the next dot 1 - supply the current dot hrf_hcu_endofline 1 Out Signal lasting 1 cycle indicating then end of the HCU read line. hrf_diurreq 1 Out Signal indicating the HCUReadLineFIFO has space for 256-bits of DIU data. hrf_diurack 1 In Acknowledge that read request has been accepted and hrf_diurreq should be de-asserted. hrf_diurdata 1 In Data from HCUReadLineFIFO to DIU. First 64-bits are bits 63:0 of 256 bit word. Second 64-bits are bits 127:64 of 256 bit word. Third 64-bits are bits 191:128 of 256 bit word. Fourth 64-bits are bits 255:192 of 256 bit word. hrf_diurvalid 1 In Signal indicating data on hrf_diurdata is valid. hrf_diuidle 1 Out Signal indicating DIU state-machine is in the IDLE state. 25.8.9.1 General Description

The HCUReadLineFIFO sub-block comprises a double 256-bit buffer between the HCU and the DIU Interface and Address Generator sub-block. The FIFO is implemented as 8 times 64-bit words. The FIFO is written by the DIU Interface and Address Generator sub-block and read by the HCU.

The DIU Interface and Address Generation (DAG) sub-block interface of the HCUReadLineFIFO is identical to the LBDPrevLineFIFO DIU interface.

Whenever 4 locations in the FIFO are free the FIFO will request 256-bits of data from the DAG sub-block by asserting hrf_diurreq. A signal hrf_diurack indicates that the request has been accepted and hrf_diurreq should be de-asserted.

The data is written to the FIFO as 64-bits on hrf_diurdata[63:0] over 4 clock cycles. The signal hrf_diurvalid indicates that the data returned on hrf_diurdata[63:0] is valid. hrf_diurvalid is used to generate the FIFO write enable, write_en, and to increment the FIFO write address, write_adr[2:0]. If the HCUReadLineFIFO still has 256-bits free then hrf_diurreq should be asserted again.

The HCUReadLineFIFO generates a signal sfu_hcu_avail to indicate that it has data available for the HCU. The HCU reads single-bit data supplied on sfu_hcu_sdata. The FIFO control logic generates a signal bit select which selects the next bit of the 64-bit FIFO word to output on sfu_hcu_sdata. The signal hcu_sfu_advdot tells the HCUReadLineFIFO to supply the next dot (hrf_xadvance=1) or the current dot (hrf_xadvance=0) on sfu_hcu_sdata according to the hrf_xadvance signal from the scaling control unit in the DAG sub-block. The HCU should not generate the hcu_sfu_advdot signal until sfu_hcu_avail is true. The HCU can therefore stall waiting for the sfu_hcu_avail signal.

When the entire current 64-bit FIFO word has been read by the HCU hcu_sfu_advdot will cause the next word to be popped from the FIFO.

The last 256-bit word for a line read from DRAM and written into the HCUReadLineFIFO can contain dots or extra padding which should not be output to the HCU. A counter in the HCUReadLineFIFO, hcuadvdot_count[15:0], counts the number of hcu_sfu_advdot strobes received from the HCU. When the count equals hcu_num_dots[15:0] the HCUReadLineFIFO must adjust the FIFO read address to point to the next 256-bit word boundary in the FIFO. This can be achieved by considering the FIFO read address, read adr[2:0], will require 3 bits to address 8 locations of 64-bits. The next 256-bit aligned address is calculated by inverting the MSB of the read_adr and setting all other bits to 0.

If (hcuadvdot_count == hcu_num_dots) then read_adr[1:0] = b00 read_adr[2] = ~read_adr[2]

The DIU Interface and Address Generator sub-block scaling unit also needs to know when hcuadvdot_count equals hcu_num_dots. This condition is exported from the HCUReadLineFIFO as the signal hrf_hcu_endofline. When the hrf_hcu_endofline is asserted the scaling unit will decide based on vertical scaling whether to go back to the start of the current line or go onto the next line.

25.8.9.2 DRAM Access Limitation

The SFU must output 1 bit/cycle to the HCU. Since HCUNumDots may not be a multiple of 256 bits the last 256-bit DRAM word on the line can contain extra zeros. In this case, the SFU may not be able to provide 1 bit/cycle to the HCU. This could lead to a stall by the SFU. This stall could then propagate if the margins being used by the HCU are not sufficient to hide it. The maximum stall can be estimated by the calculation: DRAM service period—X scale factor * dots used from last DRAM read for HCU line.

25.8.10 DIU Interface and Address Generator Sub-block

TABLE 168 DIU Interface and Address Generator Additional IO Description Port name Pins I/O Description Internal LBDPrevLineFIFO Inputs plf_diurreq 1 In Signal indicating the LBDPrevLineFIFO has 256-bits of data free. plf_diurack 1 Out Acknowledge that read request has been accepted and plf_diurreq should be de-asserted. plf_diurdata 1 Out Data from the DIU to LBDPrevLineFIFO. First 64-bits are bits 63:0 of 256 bit word Second 64-bits are bits 127:64 of 256 bit word Third 64-bits are bits 191:128 of 256 bit word Fourth 64-bits are bits 255:192 of 256 bit word plf_diurrvalid 1 Out Signal indicating data on plf_diurdata is valid. plf_diuidle 1 In Signal indicating DIU state-machine is in the IDLE state. Internal LBDNextLineFIFO Inputs nlf_diuwreq 1 In Signal indicating the LBDNextLineFIFO has 256-bits of data for writing to the DIU. nlf_diuwack 1 Out Acknowledge from DIU that write request has been accepted and write data can be output on nlf_diuwdata together with nlf_diuwvalid. nlf_diuwdata 1 In Data from LBDNextLineFIFO to DIU Interface. First 64-bits are bits 63:0 of 256 bit word Second 64-bits are bits 127:64 of 256 bit word Third 64-bits are bits 191:128 of 256 bit word Fourth 64-bits are bits 255:192 of 256 bit word nlf_diuwvalid 1 In Signal indicating that data on wlf_diuwdata is valid. Internal HCUReadLineFIFO Inputs hrf_hcu_endofline 1 In Signal lasting 1 cycle indicating then end of the HCU read line. hrf_xadvance 1 Out Signal from horizontal scaling unit 1 - supply the next dot 1 - supply the current dot hrf_diurreq 1 In Signal indicating the HCUReadLineFIFO has space for 256-bits of DIU data. hrf_diurack 1 Out Acknowledge that read request has been accepted and hrf_diurreq should be de-asserted. hrf_diurdata 1 Out Data from HCUReadLineFIFO to DIU. First 64-bits are bits 63:0 of 256 bit word Second 64-bits are bits 127:64 of 256 bit word Third 64-bits are bits 191:128 of 256 bit word Fourth 64-bits are bits 255:192 of 256 bit word hrf_diurvalid 1 Out Signal indicating data on plf_diurdata is valid. hrf_diuidle 1 In Signal indicating DIU state-machine is in the IDLE state. 25.8.10.1 General Description

The DIU Interface and Address Generator (DAG) sub-block manages the bi-level buffer in DRAM. It has a DIU Write Interface for the LBDNextLineFIFO and a DIU Read Interface shared between the HCUReadLineFIFO and LBDPrevLineFIFO.

All DRAM address management is centralised in the DAG. DRAM access is pre-emptive i.e. after a FIFO unit has made an access then as soon as the FIFO has space to read or data to write a DIU access will be requested immediately. This ensures there are no unnecessary stalls introduced e.g. at the end of an LBD or HCU line.

The control logic for horizontal and vertical non-integer scaling logic is completely contained in the DAG sub-block. The scaling control unit exports the hlf_xadvance signal to the HCUReadLineFIFO which indicates whether to replicate the current dot or supply the next dot for horizontal scaling.

25.8.10.2 DIU Write Interface

The LBDNextLineFIFO generates all the DIU write interface signals directly except for sfu_diu_wadr[21:5] which is generated by the Address Generation logic The DIU request from the LBDNextLineFIFO will be negated if its respective address pointer in DRAM is invalid i.e. nlf_adrvalid=0. The implementation must ensure that no erroneous requests occur on sfu_diu_wreq.

25.8.10.3 DIU Read Interface

Both HCUReadLineFIFO and LBDPrevLineFIFO share the read interface. If both sources request simultaneously, then the arbitration logic implements a round-robin sharing of read accesses between the HCUReadLineFIFO and LBDPrevLineFIFO.

The DIU read request arbitration logic generates a signal, select_hrfplf, which indicates whether the DIU access is from the HCUReadLineFIFO or LBDPrevLineFIFO (0=HCUReadLineFIFO, 1=LBDPrevLineFIFO). FIG. 171 shows select_hrfplf multiplexing the returned DIU acknowledge and read data to either the HCUReadLineFIFO or LBDPrevLineFIFO.

The DIU read request arbitration logic is shown in FIG. 172. The arbitration logic will select a DIU read request on hrf_diurreq or plf_diurreq and assert sfu_diu_rreq which goes to the DIU. The accompanying DIU read address is generated by the Address Generation Logic. The select signal select_hrfplf will be set according to the arbitration winner (0=HCUReadLineFIFO, 1=LBDPrevLineFIFO). sfu_diu_rreq is cleared when the DIU acknowledges the request on diu_sfu_rack. Arbitration cannot take place again until the DIU state-machine of the arbitration winner is in the idle state, indicated by diu_idle. This is necessary to ensure that the DIU read data is multiplexed back to the FIFO that requested it.

The DIU read requests from the HCUReadLineFIFO and LBDPrevLineFIFO will be negated if their respective addresses in DRAM are invalID, hrf_adrvalid=0 or plf_adrvalid=0. The implementation must ensure that no erroneous requests occur on sfu_diu_rreq.

If the HCUReadLineFIFO and LBDPrevLineFIFO request simultaneously, then if the request is not following immediately another DIU read port access, the arbitration logic will choose the HCUReadLineFIFO by default. If there are back to back requests to the DIU read port then the arbitration logic implements a round-robin sharing of read accesses between the HCUReadLineFIFO and LBDPrevLineFIFO.

A pseudo-code description of the DIU read arbitration is given below.

// history is of type {none, hrf, plf}, hrf is HCUReadLineFIFO, plf is LBDPrevLineFIFO // initialisation on reset select_hrfplf = 0 // default choose hrf history = none // no DIU read access immediately preceding // state-machine is busy between asserting sfu_diu_rreq and diu_idle = 1 // if DIU read requester state-machine is in idle state then de-assert busy if (diu_idle == 1) then busy = 0 //if acknowledge received from DIU then de-assert DIU request if (diu_sfu_rack == 1) then //de-assert request in response to acknowledge sfu_diu_rreq = 0 // if not busy then arbitrate between incoming requests // if request detected then assert busy if (busy == 0) then //if there is no request if (hrf_diurreq == 0) AND (plf_diurreq == 0) then sfu_diu_rreq = 0 history = none // else there is a request else { // assert busy and request DIU read access busy = 1 sfu_diu_rreq = 1 // arbitrate in round-robin fashion between the requestors // if only HCUReadLineFIFO requesting choose HCUReadLineFIFO if (hrf_diurreq == 1) AND (plf_diurreq == 0) then history = hrf select_hrfplf = 0 // if only LBDPrevLineFIFO requesting choose LBDPrevLineFIFO if (hrf_diurreq == 0) AND (plf_diurreq == 1) then history = plf select_hrfplf = 1 //if both HCUReadLineFIFO and LBDPrevLineFIFO requesting if (hrf_diurreq == 1) AND (plf_diurreq == 1) then // no immediately preceding request choose HCUReadLineFIFO if (history == none) then history = hrf select_hrfplf = 0 // if previous winner was HCUReadLineFIFO choose LBDPrevLineFIFO elsif (history == hrf) then history = plf select_hrfplf = 1 // if previous winner was LBDPrevLineFIFO choose HCUReadLineFIFO elsif (history == plf) then history = hrf select_hrfplf = 0 // end there is a request } 25.8.10.4 Address Generation Logic

The DIU interface generates the DRAM addresses of data read and written by the SFU's FIFOs.

A write request from the LBDNextLineFIFO on nlf_diuwreq causes a write request from the DIU Write Interface. The Address Generator supplies the DRAM write address on sfu_diu_wadr[21:5]A winning read request from the DIU read request arbitration logic causes a read request from the DIU Read Interface. The Address Generator supplies the DRAM read address on sfu_diu_radr[21:5].

The address generator is configured with the number of DRAM words to read in a HCU line, hcu_dram_words, the first DRAM address of the SFU area, start_sfu_adr[21:5], and the last DRAM address of the SFU area, end_sfu_adr[21:5].

Note hcu_dram_words configuration register specifies the the number of DRAM words consumed per line in the HCU, while lbd_dram_words specifies the number of DRAM words generated per line by the LBD. These values are not required to be the same.

For example the LBD may store 10 DRAM words per line (lbd_dram_words=10), but the HCU may consume 5 DRAM words per line. In such case the hcu_dram_words would be set to 5 and the HCU Read Line FIFO would trigger a new line after it had consumed 5 DRAM words (via hrf_hcu_endofline).

Address Generation

There are four address pointers used to manage the bi-level DRAM buffer:

-   a. hcu_readline_rd_adr is the read address in DRAM for the     HCUReadLineFIFO. -   b. hcu_startreadline_adr is the start address in DRAM for the     current line being read by the HCUReadLineFIFO. -   c. lbd_nextline_wr_adr is the write address in DRAM for the     LBDNextLineFIFO. -   d. lbd_prevline_rd_adr is the read address in DRAM for the     LBDPrevLineFIFO.

The current value of these address pointers are readable by the CPU.

Four corresponding address valid flags are required to indicate whether the address pointers are valID, based on whether the FIFOs are full or empty.

-   a. hlf_adrvalID, derived from hrf_nlf_fifo_emp -   b. hlf_start_adrvalID, derived from start_hrd_nlf_fifo_emp -   c. nlf_adrvalid. derived from nlf_plf_fifo_full and     nlf_hrf_fifo_full -   d. plf_adrvalid. derived from plf_nlf_fifo_emp

DRAM requests from the FIFOs will not be issued to the DIU until the appropriate address flag is valid.

Once a request has been acknowledged, the address generation logic can calculate the address of the next 256-bit word in DRAM, ready for the next request.

Rules for Address Pointers

The address pointers must obey certain rules which indicate whether they are valid:

-   a. hcu_readline_rd_adr is only valid if it is reading earlier in the     line than lbd_nextline_wr_adr is writing i.e. the fifo is not empty -   b. The SFU (lbd_nextline_wr_adr) cannot overwrite the current line     that the HCU is reading from (hcu_startreadline_adr) i.e. the fifo     is not full, when compared with the HCU read line pointer -   c. The LBDNextLineFIFO (lbd_nextline_wr_adr) must be writing earlier     in the line than LBD-PrevLineFIFO (lbd_prevline_rd_adr) is reading     and must not overwrite the current line that the HCU is reading from     i.e. the fifo is not full when compared to the PrevLineFifo read     pointer -   d. The LBDPrevLineFIFO (lbd_prevline_rd_adr) can read right up to     the address that LBDNextLineFIFO (lbd_nextline_wr_adr) is writing     i.e the fifo is not empty. -   e. At startup i.e. when sfu_go is asserted, the pointers are reset     to start_sfu_adr[21:5]. -   f. The address pointers can wrap around the SFU bi-level store area     in DRAM.     Address Generator Pseudo-Code:

Initialization:

if (sfu_go rising edge) then //initialise address pointers to start of SFU address space lbd_prevline_rd_adr = start_sfu_adr[21:5] lbd_nextline_wr_adr = start_sfu_adr[21:5] hcu_readline_rd_adr = start_sfu_adr[21:5] hcu_startreadline_adr = start_sfu_adr[21:5] lbd_nextline_wr_wrap = 0 lbd_prevline_rd_wrap = 0 hcu_startreadline_wrap = 0 hcu_readline_rd_wrap = 0 }

Determine FIFO fill and empty status:

// calculate which FIFOs are full and empty plf_nlf_fifo_emp = (lbd_prevline_rd_adr == lbd_nextline_wr_adr) AND (lbd_prevline_rd_wrap == lbd_nextline_wr_wrap) nlf_plf_fifo_full = (lbd_nextline_wr_adr == lbd_prevline_rd_adr) AND (lbd_prevline_rd_wrap != lbd_nextline_wr_wrap) nlf_hrf_fifo_full = (lbd_nextline_wr_adr == hcu_startreadline_adr ) AND (hcu_startreadline_wrap != lbd_nextline_wr_wrap ) // hcu start address can jump addresses and so needs comparitor if (hcu_startreadline_wrap == lbd_nextline_wr_wrap) then start_hrf_nlf_fifo_emp = (hcu_startreadline_adr >=lbd_nextline_wr_adr) else start_hrf_nlf_fifo_emp = NOT(hcu_startreadline_adr >=lbd_nextline_wr_adr) // hcu read address can jump addresses and so needs comparitor if (hcu_readline_rd_wrap == lbd_nextline_wr_wrap) then hrf_nlf_fifo_emp = (hcu_readline_rd_adr >=lbd_nextline_wr_adr) else hrf_nlf_fifo_emp = NOT(hcu_readline_rd_adr >=lbd_nextline_wr_adr)

Address ointer updating:

// LBD Next line FIFO // if DIU write acknowledge and LBDNextLineFIFO is not full with reference to PLF and HRF if (diu_sfu_wack == 1 AND nlf_plf_fifo_full != 1 AND nlf_hrf_fifo_full !=1 ) then if (lbd_nextline_wr_adr == end sfu adr) then // if end of SFU address range lbd_nextline_wr_adr = start_sfu_adr // go to start of SFU address range lbd_nextline_wr_wrap= NOT (lbd_nextline_wr_wrap) // invert the wrap bit else lbd_nextline_wr_adr++ // increment address pointer // LBD PrevLine FIFO //if DIU read acknowledge and LBDPrevLineFIFO is not empty if (diu_sfu_rack == 1 AND select_hrfplf == 1 AND plf_nlf_fifo_emp !=1) then if (lbd_prevline_rd_adr == end_sfu_adr) then lbd_prevline_rd_adr = start_sfu_adr // go to start of SFU address range lbd_prevline_rd_wrap= NOT (lbd_prevline_rd_wrap) // invert the wrap bit else lbd_prevline_rd_adr++ // increment address pointer // HCU ReadLine FIFO // if DIU read acknowledge and HCUReadLineFIFO fifo is not empty if (diu_sfu_rack == 1 AND select_hrfplf == 0 AND hrf_nlf_fifo_emp != 1) then // going to update hcu read line address if (hrf_hcu_endofline == 1) AND (hrf_yadvance == 1) then { // read the next line from DRAM // advance to start of next HCU line in DRAM hcu_startreadline_adr = hcu_startreadline_adr + lbd_dram_words offset = hcu_startreadline_adr − end_sfu_adr − 1 // allow for address wraparound if (offset >= 0) then hcu_startreadline_adr = start_sfu_adr + offset hcu_startreadline_wrap= NOT(hcu_startreadline_wrap) hcu_readline_rd_adr = hcu_startreadline_adr hcu_readline_rd_wrap= hcu_startreadline_wrap } elsif (hrf_hcu_endofline == 1) AND (hrf_yadvance == 0) then hcu_readline_rd_adr = hcu_startreadline_adr // restart and re-use the same line hcu_readline_rd_wrap= hcu_startreadline_wrap elsif (hcu readline rd adr == end sfu adr) then // check if the FIFO needs to wrap space hcu_readline_rd_adr = start_sfu_adr // go to start of SFU address space hcu_readline_rd_wrap= NOT (hcu_readline_rd_wrap) else hcu_readline_rd_adr ++ // increment address pointer 25.8.10.4.1 X Scaling of Data for HCUReadLineFIFO

The signal hcu_sfu_advdot tells the HCUReadLineFIFO to supply the next dot or the current dot on sfu_hcu_sdata according to the hrf_xadvance signal from the scaling control unit. When hrf_xadvance is 1 the HCUReadLineFIFO should supply the next dot. When hrf_xadvance is 0 the HCUReadLineFIFO should supply the current dot.

The algorithm for non-integer scaling is described in the pseudocode below. Note, x_scale_count should be loaded with x_start_count after reset and at the end of each line. The end of the line is indicated by hrf_hcu_endofline from the HCUReadLineFIFO.

if (hcu_sfu_advdot == 1) then if (x_scale_count + x_scale_denom − x_scale_num >= 0) then x_scale_count = x_scale_count + x_scale_denom − x_scale_num hrf_xadvance = 1 else x_scale_count = x_scale_count + x_scale_denom hrf_xadvance = 0 else x_scale_count = x_scale_count hrf_xadvance = 0 25.8.10.4.2 Y Scaling of Data for HCUReadLineFIFO

The HCUReadLineFIFO counts the number of hcu_sfu_advdot strobes received from the HCU. When the count equals hcu_num_dots the HCUReadLineFIFO will assert hrf_hcu_endofline for a cycle.

The algorithm for non-integer scaling is described in the pseudocode below. Note, y_scale_count should be loaded with zero after reset.

if (hrf_hcu_endofline == 1) then if (y_scale_count + y_scale_denom − y_scale_num >= 0) then y_scale_count = y_scale_count + y_scale_denom − y_scale_num hrf_yadvance = 1 else y_scale_count = y_scale_count + y_scale_denom hrf_yadvance = 0 else y_scale_count = y_scale_count hrf_yadvance = 0

When the hrf_hcu_endofline is asserted the Y scaling unit will decide whether to go back to the start of the current line, by setting hrf_yadvance=0, or go onto the next line, by setting hrf_yadvance=1.

FIG. 176 shows an overview of X and Y scaling for HCU data.

26 Tag Encoder (TE)

26.1 Overview

The Tag Encoder (TE) provides functionality for Netpage-enabled applications, and typically requires the presence of IR ink (although K ink can be used for tags in limited circumstances). The TE encodes fixed data for the page being printed, together with specific tag data values into an error-correctable encoded tag which is subsequently printed in infrared or black ink on the page. The TE places tags on a triangular grID, and can be programmed for both landscape and portrait orientations.

Basic tag structures are normally rendered at 1600 dpi, while tag data is encoded into an arbitrary number of printed dots. The TE supports integer scaling in the Y-direction while the TFU supports integer scaling in the X-direction. Thus, the TE can render tags at resolutions less than 1600 dpi which can be subsequently scaled up to 1600 dpi.

The output from the TE is buffered in the Tag FIFO Unit (TFU) which is in turn used as input by the HCU. In addition, a te_finishedband signal is output to the end of band unit once the input tag data has been loaded from DRAM. The high level data path is shown by the block diagram in FIG. 177.

After passing through the HCU, the tag plane is subsequently printed with an infrared-absorptive ink that can be read by a Netpage sensing device. Since black ink can be IR absorptive, limited functionality can be provided on offset-printed pages using black ink on otherwise blank areas of the page—for example to encode buttons. Alternatively an invisible infrared ink can be used to print the position tags over the top of a regular page. However, if invisible IR ink is used, care must be taken to ensure that any other printed information on the page is printed in infrared-transparent CMY ink, as black ink will obscure the infrared tags. The monochromatic scheme was chosen to maximize dynamic range in blurry reading environments.

When multiple SoPEC chips are used for printing the same side of a page, it is possible that a single tag will be produced by two SoPEC chips. This implies that the TE must be able to print partial tags.

The throughput requirement for the SoPEC TE is to produce tags at half the rate of the PEC1 TE.

Since the TE is reused from PEC1, the SoPEC TE over-produces by a factor of 2.

In PEC1, in order to keep up with the HCU which processes 2 dots per cycle, the tag data interface has been designed to be capable of encoding a tag in 63 cycles. This is actually accomplished in approximately 52 cycles within PEC1. If the SoPEC TE were to be modified from two dots production per cycle to a nominal one dot per cycle it should not lose the 63/52 cycle performance edge attained in the PEC1 TE.

26.2 What are Tags?

The first barcode was described in the late 1940's by Woodland and Silver, and finally patented in 1952 (U.S. Pat. No. 2,612,994) when electronic parts were scarce and very expensive. Now however, with the advent of cheap and readily available computer technology, nearly every item purchased from a shop contains a barcode of some description on the packaging. From books to CDs, to grocery items, the barcode provides a convenient way of identifying an object by a product number. The exact interpretation of the product number depends on the type of barcode. Warehouse inventory tracking systems let users define their own product number ranges, while inventory in shops must be more universally encoded so that products from one company don't overlap with products from another company. Universal Product Codes (UPC) were introduced in the mid 1970's at the request of the National Association of Food Chains for this very reason. Barcodes themselves have been specified in a large number of formats. The older barcode formats contain characters that are displayed in the form of lines. The combination of black and white lines describe the information the barcodes contains. Often there are two types of lines to form the complete barcode: the characters (the information itself) and lines to separate blocks for better optical recognition. While the information may change from barcode to barcode, the lines to separate blocks stays constant. The lines to separate blocks can therefore be thought of as part of the constant structural components of the barcode.

Barcodes are read with specialized reading devices that then pass the extracted data onto the computer for further processing. For example, a point-of-sale scanning device allows the sales assistant to add the scanned item to the current sale, places the name of the item and the price on a display device for verification etc. Light-pens, gun readers, scanners, slot readers, and cameras are among the many devices used to read the barcodes.

To help ensure that the data extracted was read correctly, checksums were introduced as a crude form of error detection. More recent barcode formats, such as the Aztec 2D barcode developed by Andy Longacre in 1995 (U.S. Pat. No. 5,591,956), but now released to the public domain, use redundancy encoding schemes such as Reed-Solomon. Reed Solomon encoding is adequately discussed in [28], [30] and [34]. The reader is advised to refer to these sources for background information. Very often the degree of redundancy encoding is user selectable.

More recently there has also been a move from the simple one dimensional barcodes (line based) to two dimensional barcodes. Instead of storing the information as a series of lines, where the data can be extracted from a single dimension, the information is encoded in two dimensions. Just as with the original barcodes, the 2D barcode contains both information and structural components for better optical recognition. FIG. 178 shows an example of a QR Code (Quick Response Code), developed by Denso of Japan (U.S. Pat. No. 5,726,435). Note the barcode cell is comprised of two areas: a data area (depends on the data being stored in the barcode), and a constant position detection pattern. The constant position detection pattern is used by the reader to help locate the cell itself, then to locate the cell boundaries, to allow the reader to determine the original orientation of the cell (orientation can be determined by the fact that there is no 4th corner pattern).

The number of barcode encoding schemes grows daily. Yet very often the hardware for producing these barcodes is specific to the particular barcode format. As printers become more and more embedded, there is an increasing desire for real-time printing of these barcodes. In particular, Netpage enabled applications require the printing of 2D barcodes (or tags) over the page, preferably in infra-red ink. The tag encoder in SoPEC uses a generic barcode format encoding scheme which is particularly suited to real-time printing. Since the barcode encoding format is generic, the same rendering hardware engine can be used to produce a wide variety of barcode formats.

Unfortunately the term “barcode” is interpreted in different ways by different people. Sometimes it refers only to the data area component, and does not include the constant position detection pattern. In other cases it refers to both data and constant position detection pattern.

We therefore use the term tag to refer to the combination of data and any other components (such as position detection pattern, blank space etc. surround) that must be rendered to help hold or locate/read the data. A tag therefore contains the following components:

-   -   data area(s). The data area is the whole reason that the tag         exists. The tag data area(s) contains the encoded data         (optionally redundancy-encoded, perhaps simply checksummed)         where the bits of the data are placed within the data area at         locations specified by the tag encoding scheme.     -   constant background patterns, which typically includes a         constant position detection pattern. These help the tag reader         to locate the tag. They include components that are easy to         locate and may contain orientation and perspective information         in the case of 2D tags. Constant background patterns may also         include such patterns as a blank area surrounding the data area         or position detection pattern. These blank patterns can aid in         the decoding of the data by ensuring that there is no         interference between tags or data areas.

In most tag encoding schemes there is at least some constant background pattern, but it is not necessarily required by all. For example, if the tag data area is enclosed by a physical space and the reading means uses a non-optical location mechanism (e.g. physical alignment of surface to data reader) then a position detection pattern is not required.

Different tag encoding schemes have different sized tags, and have different allocation of physical tag area to constant position detection pattern and data area. For example, the QR code has 3 fixed blocks at the edges of the tag for position detection pattern (see FIG. 178) and a data area in the remainder. By contrast, the Netpage tag structure (see FIGS. 179 and 180) contains a circular locator component, an orientation feature, and several data areas. FIG. 179( a) shows the Netpage tag constant background pattern in a resolution independent form. FIG. 179( b) is the same as FIG. 179( a), but with the addition of the data areas to the Netpage tag. FIG. 180 is an example of dot placement and rendering to 1600 dpi for a Netpage tag. Note that in FIG. 180 a single bit of data is represented by many physical output dots to form a block within the data area.

26.2.1 Contents of the Data Area

The data area contains the data for the tag.

Depending on the tag's encoding format, a single bit of data may be represented by a number of physical printed dots. The exact number of dots will depend on the output resolution and the target reading/scanning resolution. For example, in the QR code (see FIG. 178), a single bit is represented by a dark module or a light module, where the exact number of dots in the dark module or light module depends on the rendering resolution and target reading/scanning resolution. For example, a dark module may be represented by a square block of printed dots (all on for binary 1, or all off for binary 0), as shown in FIG. 181.

The point to note here is that a single bit of data may be represented in the printed tag by an arbitrary printed shape. The smallest shape is a single printed dot, while the largest shape is theoretically the whole tag itself, for example a giant macrodot comprised of many printed dots in both dimensions.

An ideal generic tag definition structure allows the generation of an arbitrary printed shape from each bit of data.

26.2.2 What do the Bits Represent?

Given an original number of bits of data, and the desire to place those bits into a printed tag for subsequent retrieval via a reading/scanning mechanism, the original number of bits can either be placed directly into the tag, or they can be redundancy-encoded in some way. The exact form of redundancy encoding will depend on the tag format.

The placement of data bits within the data area of the tag is directly related to the redundancy mechanism employed in the encoding scheme. The idea is generally to place data bits together in 2D so that burst errors are averaged out over the tag data, thus typically being correctable. For example, all the bits of Reed-Solomon codeword would be spread out over the entire tag data area so to minimize being affected by a burst error.

Since the data encoding scheme and shape and size of the tag data area are closely linked, it is desirable to have a generic tag format structure. This allows the same data structure and rendering embodiment to be used to render a variety of tag formats.

26.2.2.1 Fixed and Variable Data Components

In many cases, the tag data can be reasonably divided into fixed and variable components. For example, if a tag holds N bits of data, some of these bits may be fixed for all tags while some may vary from tag to tag.

For example, the Universal product code allows a country code and a company code. Since these bits don't change from tag to tag, these bits can be defined as fixed, and don't need to be provided to the tag encoder each time, thereby reducing the bandwidth when producing many tags.

Another example is Netpage tags. A single printed page contains a number of Netpage tags. The page-id will be constant across all the tags, even though the remainder of the data within each tag may be different for each tag. By reducing the amount of variable data being passed to SoPEC's tag encoder for each tag, the overall bandwidth can be reduced.

Depending on the embodiment of the tag encoder, these parameters will be either implicit or explicit, and may limit the size of tags renderable by the system. For example, a software tag encoder may be completely variable, while a hardware tag encoder such as SoPEC's tag encoder may have a maximum number of tag data bits.

26.2.2.2 Redundancy-Encode the Tag Data within the Tag Encoder

Instead of accepting the complete number of TagData bits encoded by an external encoder, the tag encoder accepts the basic non-redundancy-encoded data bits and encodes them as required for each tag. This leads to significant savings of bandwidth and on-chip storage.

In SoPEC's case for Netpage tags, only 120 bits of original data are provided per tag, and the tag encoder encodes these 120 bits into 360 bits. By having the redundancy encoder on board the tag encoder the effective bandwidth and internal storage required is reduced to only 33% of what would be required if the encoded data was read directly.

26.3 Placement of Tags on a Page

The TE places tags on the page in a triangular grid arrangement as shown in FIG. 182.

The triangular mesh of tags combined with the restriction of no overlap of columns or rows of tags means that the process of tag placement is greatly simplified. For a given line of dots, all the tags on that line correspond to the same part of the general tag structure. The triangular placement can be considered as alternative lines of tags, where one line of tags is inset by one amount in the dot dimension, and the other line of dots is inset by a different amount. The dot inter-tag gap is the same in both lines of tag, and is different from the line inter-tag gap.

Note also that as long as the tags themselves can be rotated, portrait and landscape printing are essentially the same—the placement parameters of line and dot are swapped, but the placement mechanism is the same.

The general case for placement of tags therefore relies on a number of parameters, as shown in FIG. 183.

The parameters are more formally described in Table 169. Note that these are placement parameters and not registers.

TABLE 169 Tag placement parameters parameter description restrictions Tag height The number of dot lines in minimum 1 a tag's bounding box Tag width The number of dots in a minimum 1 single line of the tag's bounding box. The number of dots in the tag itself may vary depending on the shape of the tag, but the number of dots in the bounding box will be constant (by definition). Dot inter- The number of dots from the minimum = 0 tag gap edge of one tag's bounding box to the start of the next tag's bounding box, in the dot direction. Line inter- The number of dot lines minimum = 0 tag gap from the edge of one tag's bounding box to the start of the next tag's bounding box, in the line direction. Start Defines the status of the — Position top left dot on the page - is an offset in dot & row within the tag or the inter-tag gap. AltTagLinePosition Defines the status for the — start of the alternate row of tags. Is an offset in dot within the tag or within the dot inter-tag gap (the row position is always 0). 26.4 Basic Tag Encoding Parameters

SoPEC's tag encoder imposes range restrictions on tag encoding parameters as a direct result of on-chip buffer sizes. Table 170 lists the basic encoding parameters as well as range restrictions where appropriate. Although the restrictions were chosen to take the most likely encoding scenarIOs into account, it is a simple matter to adjust the buffer sizes and corresponding addressing to allow arbitrary encoding parameters in future implementations.

TABLE 170 Encoding parameters name definition maximum value imposed by TE W page width 2¹⁴ dotpairs or 20.48 inches at 1600 dpi S tag size typical tag size is 2 mm × 2 mm maximum tag size is 384 dots × 384 dots before scaling i.e. 6 mm × 6 mm at 1600 dpi N number of dots in each 384 dots before scaling dimension of the tag E redundancy encoding for Reed-Solomon GF(2⁴) at tag data 5:10 or 7:8 D_(F) size of fixed data 40 or 56 bits (unencoded) R_(F) size of redundancy- 120 bits encoded fixed data D_(V) size of variable data 120 or 112 bits (unencoded) R_(V) size of redundancy- 360 or 240 bits encoded variable data T tags per page width 256

The fixed data for the tags on a page need only be supplied to the TE once. It can be supplied as 40 or 56 bits of unencoded data and encoded within the TE as described in Section 26.4.1. Alternatively it can be supplied as 120 bits of pre-encoded data (encoded arbitrarily).

The variable data for the tags on a page are those 112 or 120 data bits that are variable for each tag. Variable tag data is supplied as part of the band data, and is always encoded by the TE as described in Section 26.4.1, but may itself be arbitrarily pre-encoded.

26.4.1 Redundancy Encoding

The mapping of data bits (both fixed and variable) to redundancy encoded bits relies heavily on the method of redundancy encoding employed. Reed-Solomon encoding was chosen for its ability to deal with burst errors and effectively detect and correct errors using a minimum of redundancy. Reed Solomon encoding is adequately discussed in [28], [30] and [34]. The reader is advised to refer to these sources for background information.

In this implementation of the TE we use Reed-Solomon encoding over the Galois Field GF(2⁴). Symbol size is 4 bits. Each codeword contains 15 4-bit symbols for a codeword length of 60 bits.

The primitive polynomial is p(x)=x⁴+x+1, and the generator polynomial is g(x)=(x+α)(x+α²) . . . (x+α^(2t)), where t=the number of symbols that can be corrected.

Of the 15 symbols, there are two possibilities for encoding:

-   -   RS(15, 5): 5 symbols original data (20 bits), and 10 redundancy         symbols (40 bits). The 10 redundancy symbols mean that we can         correct up to 5 symbols in error. The generator polynomial is         therefore g(x)=(x+α)(x+α²) . . . (x+α¹⁰).     -   RS(15, 7): 7 symbols original data (28 bits), and 8 redundancy         symbols (32 bits). The 8 redundancy symbols mean that we can         correct up to 4 symbols in error. The generator polynomial is         g(x)=(x+α)(x+α²) . . . (x+α⁸).

In the first case, with 5 symbols of original data, the total amount of original data per tag is 160 bits (40 fixed, 120 variable). This is redundancy encoded to give a total amount of 480 bits (120 fixed, 360 variable) as follows:

-   -   Each tag contains up to 40 bits of fixed original data.         Therefore 2 codewords are required for the fixed data, giving a         total encoded data size of 120 bits. Note that this fixed data         only needs to be encoded once per page.     -   Each tag contains up to 120 bits of variable original data.         Therefore 6 codewords are required for the variable data, giving         a total encoded data size of 360 bits.

In the second case, with 7 symbols of original data, the total amount of original data per tag is 168 bits (56 fixed, 112 variable). This is redundancy encoded to give a total amount of 360 bits (120 fixed, 240 variable) as follows:

-   -   Each tag contains up to 56 bits of fixed original data.         Therefore 2 codewords are required for the fixed data, giving a         total encoded data size of 120 bits. Note that this fixed data         only needs to be encoded once per page.     -   Each tag contains up to 112 bits of variable original data.         Therefore 4 codewords are required for the variable data, giving         a total encoded data size of 240 bits.

The choice of data to redundancy ratio depends on the application.

26.5 Data Structures Used by Tag Encoder

26.5.1 Tag Format Structure

The Tag Format Structure (TFS) is the template used to render tags, optimized so that the tag can be rendered in real time. The TFS contains an entry for each dot position within the tag's bounding box. Each entry specifies whether the dot is part of the constant background pattern or part of the tag's data component (both fixed and variable).

The TFS is very similar to a bitmap in that it contains one entry for each dot position of the tag's bounding box. The TFS therefore has TagHeight×TagWidth entries, where TagHeight matches the height of the bounding box for the tag in the line dimension, and TagWidth matches the width of the bounding box for the tag in the dot dimension. A single line of TFS entries for a tag is known as a tag line structure.

The TFS consists of TagHeight number of tag line structures, one for each 1600 dpi line in the tag's bounding box. Each tag line structure contains three contiguous tables, known as tables A, B, and C. Table A contains 384 2-bit entries, one entry for each of the maximum number of dots in a single line of a tag (see Table). The actual number of entries used should match the size of the bounding box for the tag in the dot dimension, but all 384 entries must be present. Table B contains 32 9-bit data addresses that refer to (in order of appearance) the data dots present in the particular line. All 32 entries must be present, even if fewer are used. Table C contains two 5-bit pointers into table B, and therefore comprises 10 bits. Padding of 214 bits is added. The total length of each tag line structure is therefore 5×256-bit DRAM words. Thus a TFS containing TagHeight tag line structures requires a TagHeight*160 bytes. The structure of a TFS is shown in FIG. 184.

A full description of the interpretation and usage of Tables A, B and C is given in section 26.8.3 on page 444.

26.5.1.1 Scaling a Tag

If the size of the printed dots is too small, then the tag can be scaled in one of several ways. Either the tag itself can be scaled by N dots in each dimension, which increases the number of entries in the TFS. As an alternative, the output from the TE can be scaled up by pixel replication via a scale factor greater than 1 in the both the TE and TFU.

For example, if the original TFS was 21×21 entries, and the scaling were a simple 2×2 dots for each of the original dots, we could increase the TFS to be 42×42. To generate the new TFS from the old, we would repeat each entry across each line of the TFS, and then we would repeat each line of the TFS. The net number of entries in the TFS would be increased fourfold (2×2).

The TFS allows the creation of macrodots instead of simple scaling. Looking at FIG. 185 for a simple example of a 3×3 dot tag, we may want to produce a physically large printed form of the tag, where each of the original dots was represented by 7×7 printed dots. If we simply performed replication by 7 in each dimension of the original TFS, either by increasing the size of the TFS by 7 in each dimension or putting a scale-up on the output of the tag generator output, then we would have 9 sets of 7×7 square blocks. Instead, we can replace each of the original dots in the TFS by a 7×7 dot definition of a rounded dot. FIG. 186 shows the results.

Consequently, the higher the resolution of the TFS the more printed dots can be printed for each macrodot, where a macrodot represents a single data bit of the tag. The more dots that are available to produce a macrodot, the more complex the pattern of the macrodot can be. As an example, Figure n page 461 on page Error! Bookmark not defined. shows the Netpage tag structure rendered such that the data bits are represented by an average of 8 dots×8 dots (at 1600 dpi), but the actual shape structure of a dot is not square. This allows the printed Netpage tag to be subsequently read at any orientation.

26.5.2 Raw Tag Data

The TE requires a band of unencoded variable tag data if variable data is to be included in the tag bit-plane. A band of unencoded variable tag data is a set of contiguous unencoded tag data records, in order of encounter top left of printed band from top left to lower right.

An unencoded tag data record is 128 bits arranged as follows: bits 0-111 or 0-119 are the bits of raw tag data, bit 120 is a flag used by the TE (TagIsPrinted), and the remaining 7 bits are reserved (and should be 0). Having a record size of 128 bits simplifies the tag data access since the data of two tags fits into a 256-bit DRAM word. It also means that the flags can be stored apart from the tag data, thus keeping the raw tag data completely unrestricted. If there is an odd number of tags in line then the last DRAM read will contain a tag in the first 128 bits and padding in the final 128 bits.

The TagIsPrinted flag allows the effective specification of a tag resolution mask over the page. For each tag position the TagIsPrinted flag determines whether any of the tag is printed or not. This allows arbitrary placement of tags on the page. For example, tags may only be printed over particular active areas of a page. The TagIsPrinted flag allows only those tags to be printed. TagIsPrinted is a 1 bit flag with values as shown in Table 171.

TABLE 171 TagIsPrinted values Value description 0 Don't print the tag in this tag position. Output 0 for each dot within the tag bounding box. 1 Print the tag as specified by the various tag structures. 26.5.3 DRAM Storage Requirements

The total DRAM storage required by a single band of raw tag data depends on the number of tags present in that band. Each tag requires 128 bits. Consequently if there are N tags in the band, the size in DRAM is 16N bytes.

The maximum size of a line of tags is 163×128 bits. When maximally packed, a row of tags contains 163 tags (see Table) and extends over a minimum of 126 print lines. This equates to 282 KBytes over a Letter page.

The total DRAM storage required by a single TFS is TagHeight/7 KBytes (including padding). Since the likely maximum value for TagHeight is 384 (given that SoPEC restricts TagWidth to 384), the maximum size in DRAM for a TFS is 55 KBytes.

26.5.4 DRAM Access Requirements

The TE has two separate read interfaces to DRAM for raw tag data, TD, and tag format structure, TFS.

The memory usage requirements are shown in Table 172. Raw tag data is stored in the compressed page store

TABLE 172 Memory usage requirements Block Size Description Compressed 2048 Kbytes Compressed data page store for page Bi-level, contone and raw tag store data. Tag Format 55 Kbyte (384 55 kB in PEC1 for 384 dot line Structure dot line tags (the benchmark) at 1600 dpi tags @ 2.5 mm tags 1600 dpi) ( 1/10th inch) @ 1600 dpi require 160 dot lines = 160/384 × 55 or 23 kB 2.5 mm tags @ 800 dpi require 80/384 × 55 = 12 kB

The TD interface will read 256-bits from DRAM at a time. Each 256-bit read returns 2 times 128-bit tags. The TD interface to the DIU will be a 256-bit double buffer. If there is an odd number of tags in line then the last DRAM read will contain a tag in the first 128 bits and padding in the final 128 bits.

The TFS interface will also read 256-bits from DRAM at a time. The TFS required for a line is 136 bytes. A total of 5 times 256-bit DRAM reads is required to read the TFS for a line with 192 unused bits in the fifth 256-bit word. A 136-byte double-line buffer will be implemented to store the TFS data.

The TE's DIU bandwidth requirements are summarized in Table 173.

TABLE 173 DRAM bandwidth requirements Peak Average Maximum number of Bandwidth Bandwidth Block cycles between each (bits/ (bits/ Name Direction 256-bit DRAM access cycle) cycle) TD Read Single 256 bit reads1. 1.02 1.02 TFS Read Single 256 bit reads2. 0.093 0.093 TFS is 136 bytes. This means there is unused data in the fifth 256 bit read. A total of 5 reads is required. 1: Each 2 mm tag lasts 126 dot cycles and requires 128 bits. This is a rate of 256 bits every 252 cycles. 2: 17×64 bit reads per line in PEC1 is 5×256 bit reads per line in SoPEC with unused bits in the last 256-bit read. 26.5.5 TD and TFS Bandstore Wrapping

TABLE 174 Bandstore Inputs from CDU Port Name Pins I/O Description cdu_endofbandstore[21:5] 17 In Address of the end of the current band of data. 256-bit word aligned DRAM address. cdu_startofbandstore[21:5] 17 In Address of the start of the current band of data. 256-bit word aligned DRAM address.

Both TD and TFS storage in DRAM can wrap around the bandstore area. The bounds of the band store are described by inputs from the CDU shown in Table 174. The TD and TFS DRAM interfaces therefore support bandstore wrapping. If the TD or TFS DRAM interface increments an address it is checked to see if it matches the end of bandstore address. If so, then the address is mapped to the start of the bandstore.

26.5.6 Tag Sizes

SoPEC allows for tags to be between 0 to 384 dots. A typical 2 mm tag requires 126 dots. Short tags do not change the internal bandwidth or throughput behaviours at all. Tag height is specified so as to allow the DRAM storage for raw tag data to be specified. Minimum tag width is a condition imposed by throughput limitations, so if the width is too small TE cannot consistently produce 2 dots per cycle across several tags (also there are raw tag data bandwidth implications). Thinner tags still work, they just take longer and/or need scaling.

26.6 Implementation

26.6.1 Tag Encoder Architecture

A block diagram of the TE can be seen below.

The TE writes lines of bi-level tag plane data to the TFU for later reading by the HCU. The TE is responsible for merging the encoded tag data with the tag structure (interpreted from the TFS). Y-integer scaling of tags is performed in the TE with X-integer scaling of the tags performed in the TFU. The encoded tag layer is generated 2 bits at a time and output to the TFU at this rate. The HCU however only consumes 1 bit per cycle from the TFU. The TE must provide support for 126 dot Tags (2 mm densely packed) with 108 Tags per line with 128 bits per tag.

The tag encoder consists of a TFS interface that loads and decodes TFS entries, a tag data interface that loads tag raw data, encodes it, and provides bit values on request, and a state machine to generate appropriate addressing and control signals. The TE has two separate read interfaces to DRAM for raw tag data, TD, and tag format structure, TFS.

It is possible that the raw tag data interface, the TD, to the DIU could be replaced by a hardware state machine at a later stage. This would allow flexibility in the generation of tags. Support for Y scaling needs to be added to the PEC1 TE. The PEC1 TE already allows stalling at its output during a line when tfu_te_oktowrite is deasserted.

26.6.2 Y-Scaling Output Lines

In order to support scaling in the Y direction the following modifications to the PEC1 TE are suggested to the Tag Data Interface, Tag Format Structure Interface and TE Top Level:

-   -   for Tag Data Interface: program the configuration registers of         Table, firstTagLineHeight and tagMaxLine with true value i.e.         not multiplied up by the scale factor YScale. Within the Tag         Data interface there are two counters, countx and county that         have a direct bearing on the rawTagDataAddr generation. countx         decrements as tags are read from DRAM. It is reset to         NumTags[RtdTagSense] at start of each line of tags. county is         decremented as each line of tags is completely read from DRAM         i.e. countx=0. Scaling may be performed by counting the number         of times countx reaches zero and only decrementing county when         this number reaches YScale. This will cause the TagData         Interface to read each line of tag data NumTags[RtdTagSense]*         YScale times.     -   for Tag Format Structure Interface: The implication of Y-scaling         for the TFS is that each Tag Line Structure is used YScale         times. This may be accomplished in either of two ways:     -   For each Tag Line Structure read it once from DRAM and reuse         YScale times. This involves gating the control of TFS buffer         flipping with YScale. Because of the way in which this         advTfsLine and advTagLine related functionality is coded in the         PEC1 TFS this solution is judged to be error-prone.     -   Fetch each TagLineStructure YScale times. This solution involves         controlling the activity of currTfsAddr with YScale.         -   In SoPEC the TFS must supply five addresses to the DIU to             read each individual Tag Line Structure. The DIU returns             4*64-bit words for each of the 5 accesses. This is different             from the behaviour in PEC1, where one address is given and             17 data-words were returned by the DIU.         -   Since the behaviour of the currTfsAddr must be changed to             meet the requirements of the SoPEC DIU it makes sense to             include the Y-Scaling into this change i.e. a count of the             number of completed sets of 5 accesses to the DIU is             compared to YScale. Only when this count equals YScale can             currTfsAddr be loaded with the base address of the next             lines Tag Line Structure in DRAM, otherwise it is re-loaded             with the base address of the current lines Tag Line             Structure in DRAM.     -   For Top Level: The Top Level of the TE has a counter, LinePos,         which is used to count the number of completed output lines when         in a tag gap or in a line of tags. At the start (i.e. top-left         hand dot-pair) of a gap or tag LinePos is loaded with either         TagGapLine or TagMaxLine. The value of LinePos is decremented at         last dot-pair in line. Y-Scaling may be accomplished by gating         the decrement of LinePos based on YScale value         26.6.3 TE Physical Hierarchy

FIG. 188 above illustrates the structural hierarchy of the TE. The top level contains the Tag Data Interface (TDI), Tag Format Structure (TFS), and an FSM to control the generation of dot pairs along with a clocked process to carry out the PCU read/write decoding. There is also some additional logic for muxing the output data and generating other control signals.

At the highest level, the TE state machine processes the output lines of a page one line at a time, with the starting position either in an inter-tag gap or in a tag (a SoPEC may be only printing part of a tag due to multiple SoPECs printing a single line).

If the current position is within an inter-tag gap, an output of 0 is generated. If the current position is within a tag, the tag format structure is used to determine the value of the output dot, using the appropriate encoded data bit from the fixed or variable data buffers as necessary. The TE then advances along the line of dots, moving through tags and inter-tag gaps according to the tag placement parameters.

26.6.4 IO Definitions

TABLE 175 TE Port List Port Name Pins I/O Description Clocks and Resets Pclk 1 In SoPEC Functional clock. prst_n 1 In Global reset signal. Bandstore Signals cdu_endofbandstore[21:5] 17 In Address of the end of the current band of data. 256-bit word aligned DRAM address. cdu_startofbandstore[21:5] 17 In Address of the start of the current band of data. 256-bit word aligned DRAM address. te_finishedband 1 Out TE finished band sig- nal to PCU and ICU. PCU Interface data and control signals pcu_addr[8:2] 7 In PCU address bus. 7 bits are required to decode the address space for this block. pcu_dataout[31:0] 32 In Shared write data bus from the PCU. te_pcu_datain[31:0] 32 Out Read data bus from the TE to the PCU. pcu_rwn 1 In Common read/not-write signal from the PCU. pcu_te_sel 1 In Block select from the PCU. When pcu_te_sel is high both pcu_addr and pcu_dataout are valid. te_pcu_rdy 1 Out Ready signal to the PCU. When te_pcu_rdy is high it indicates the last cycle of the access. For a write cycle this means pcu_dataout has been registered by the block and for a read cycle this means the data on te_pcu_datain is valid. TD (raw Tag Data) DIU Read Interface signals td_diu_rreq 1 Out TD requests DRAM read. A read request must be accompanied by a valid read address. td_diu_radr[21:5] 17 Out TD read address to DIU. 17 bits wide (256-bit aligned word). diu_td_rack 1 In Acknowledge from DIU that TD read request has been accepted and new read address can be placed on te_diu_radr. diu_data[63:0] 64 In Data from DIU to TE. First 64-bits are bits 63:0 of 256 bit word; Second 64-bits are bits 127:64 of 256 bit word; Third 64-bits are bits 191:128 of 256 bit word; Fourth 64-bits are bits 255:192 of 256 bit word. diu_td_rvalid 1 In Signal from DIU telling TD that valid read data is on the diu_data bus. TFS (Tag Format Structure) DIU Read Interface signals tfs_diu_rreq 1 Out TFS requests DRAM read. A read request must be accompanied by a valid read address. tfs_diu_radr[21:5] 17 Out TFS Read address to DIU 17 bits wide (256-bit aligned word). diu_tfs_rack 1 In Acknowledge from DIU that TFS read request has been accepted and new read address can be placed on tfs_diu_radr. diu_data[63:0] 64 In Data from DIU to TE. First 64-bits are bits 63:0 of 256 bit word; Second 64-bits are bits 127:64 of 256 bit word; Third 64-bits are bits 191:128 of 256 bit word; Fourth 64-bits are bits 255:192 of 256 bit word. diu_tfs_rvalid 1 In Signal from DIU telling TFS that valid read data is on the diu_data bus. TFU Interface data and control signals tfu_te_oktowrite 1 In Ready signal indicating TFU has space available and is ready to be written to. Also asserted from the point that the TFU has recieved its expected number of bytes for a line until the next te_tfu_wradvline te_tfu_wdata[7:0] 8 Out Write data for TFU. te_tfu_wdatavalid 1 Out Write data valid signal. This signal remains high whenever there is valid output data on te_tfu_wdata te_tfu_wradvline 1 Out Advance line signal strobed when the last byte in a line is placed on te_tfu_wdata 26.6.5 Configuration Registers

The configuration registers in the TE are programmed via the PCU interface. Refer to section 21.8.2 on page 321 for the description of the protocol and timing diagrams for reading and writing registers in the TE. Note that since addresses in SoPEC are byte aligned and the PCU only supports 32-bit register reads and writes the lower 2 bits of the PCU address bus are not required to decode the address space for the TE. Table 176 lists the configuration registers in the TE. Registers which address DRAM are 64-bit DRAM word aligned as this is the case for the PEC1 TE. SoPEC assumes a 256-bit DRAM word size. If the TE can be easily modified then the DRAM word addressing should be modified to 256-bit word aligned addressing. Otherwise, software should program these the 64-bit word aligned addresses on a 256-bit DRAM word boundary.

TABLE 176 TE Configuration Registers value Address register on TE_base+ name #bits reset description Control registers 0x00 Reset 1 1 A write to this register causes a reset of the TE. This register can be read to indicate the reset state: 0 - reset in progress 1 - reset not in progress 0x04 Go 1 0 Writing 1 to this register starts the TE. Writing 0 to this register halts the TE. When Go is deasserted the state-machines go to their idle states but all counters and configuration registers keep their values. When Go is asserted all counters are reset, but configuration registers keep their values (i.e. they don't get reset). NextBandEnable is cleared when Go is asserted. The TFU must be started before the TE is started. This register can be read to determine if the TE is running (1 = running, 0 = stopped). Setup registers (constant for processing of a page) 0x40 TfsStartAdr 19 0 Points to the first word of the (64-bit aligned first TFS line in DRAM. DRAM address - should start at a 256-bit aligned location) 0x44 TfsEndAdr 19 0 Points to the first word of the (64-bit aligned last TFS line in DRAM. DRAM address - should start at a 256-bit aligned location) 0x48 TfsFirstLineAdr 19 0 Points to the first word of the (64-bit aligned first TFS line to be DRAM address) encountered on the page. If the start of the page is in an inter-tag gap, then this value will be the same as TFSStartAdr since the first tag line reached will be the top line of a tag. 0x4C DataRedun 1 0 Defines the data to redundancy ratio for the Reed Solomon encoder. Symbol size is always 4 bits, Code- word size is always 15 symbols (60 bits). 0-5 data symbols (20 bits), 10 redundancy symbols (40 bits) 1-7 data symbols (28 bits), 8 redundancy symbols (32 bits) 0x50 Decode2DEn 1 0 Determines whether or not the data bits are to be 2D decoded rather than redundancy encoded (each 2 bits of the data bits becomes 4 output data bits). 0 = redundancy encode data 1 = decode each 2 bits of data into 4 bits 0x54 VariableDataPresent 1 0 Defines whether or not there is variable data in the tags. If there is none, no attempt is made to read tag data, and tag encoding should only reference fixed tag data. 0x58 EncodeFixed 1 0 Determines whether or not the lower 40 (or 56) bits of fixed data should be encoded into 120 bits or simply used as is. 0x5C TagMaxDotpairs 8 0 The width of a tag in dot- pairs, minus 1. Minimum 0, Maximum = 191. 0x60 TagMaxLine 9 0 The number of lines in a tag, minus 1. Minimum 0, Maximum = 383. 0x64 TagGapDot 14 0 The number of dot pairs between tags in the dot dimension minus 1. Only valid if TagGapPresent[bit 0] = 1. 0x68 TagGapLine 14 0 Defines the number of dotlines between tags in the line dimension minus 1. Only valid if TagGapPresent[bit1] = 1. 0x6C DotPairsPerLine 14 0 Number of output dot pairs to generate per tag line. 0x70 DotStartTagSense 2 0 Determines for the first/even (bit 0) and second/odd (bit 1) rows of tags whether or not the first dot position of the line is in a tag. 1 = in a tag, 0 = in an inter-tag gap. 0x74 TagGapPresent 2 0 Bit 0 is 1 if there is an inter-tag gap in the dot dimension, and 0 if tags are tightly packed. Bit 1 is 1 if there is an inter-tag gap in the line dimension, and 0 if tags are tightly packed. 0x78 YScale 8 1 Tag scale factor in Y direction. Output lines to the TFU will be generated YScale times. 0x80 to DotStartPos 2 × 14 0 Determines for the first/even 0x84 (0) and second/odd (1) rows of tags the number of dotpairs remaining minus 1, in either the tag or inter-tag gap at the start of the line. 0x88 to NumTags 2 × 8  0 Determines for the first/even 0x8C and second/odd rows of tags how many tags are present in a line (equals number of tags minus 1). Setup band related registers 0xC0 NextBandStartTagDataAdr Holds the value of (64-bit aligned StartTagDataAdr for the next DRAM address - band. This value is copied to should start at StartTagDataAdr when a 256-bit DoneBand is 1 and aligned location) NextBandEnable is 1, or when Go transitions from 0 to 1. 0xC4 NextBandEndOfTagData Holds the value of (64-bit aligned EndOfTagData for the next DRAM address) band. This value is copied to EndOfTagData when DoneBand is 1 and NextBandEnable is 1, or when Go transitions from 0 to 1. 0xC8 NextBandFirstTagLineHeight 9 0 Holds the value of FirstTagLineHeight for the next band. This value is copied to FirstTagLineHeight when DoneBand gets is 1 and NextBandEnable is 1, or when Go transitions from 0 to 1. 0xCC NextBandEnable When NextBandEnable is 1 and DoneBand is 1, then when te_finishedband is set at the end of a band: NextBandStartTagDataAdr is copied to StartTagDataAdr NextBandEndOfTagData is copied to EndOfTagData NextBandFirstTagLineHeight is copied to FirstTagLineHeight DoneBand is cleared NextBandEnable is cleared. NextBandEnable is cleared when Go is asserted. Read-only band related registers 0xD0 DoneBand 1 0 Specifies whether the tag data interface has finished loading all the tag data for the band. It is cleared to 0 when Go transitions from 0 to 1. When the tag data interface has finished loading all the tag data for the band, the te_finishedband signal is given out and the DoneBand flag is set. If NextBandEnable is 1 at this time then startTagDataAdr, endOfTagData and firstTaglineHeight are updated with the values for the next band and DoneBand is cleared. Processing of the next band starts immediately. If NextBandEnable is 0 then the remainder of the TE will continue to run,, while the read control unit waits for NextBandEnable to be set before it restarts. Read only. 0xD4 StartTagDataAdr 19 0 The start address of the (64-bit aligned current row of raw tag data. DRAM address - This is initially points to the should start at first word of the band's tag a 256-bit data, which should be aligned aligned location) to a 128-bit boundary (i.e. the lower bit of this address should be 0). Read only. 0xD8 EndOfTagData 19 0 Points to the address of the (64-bit aligned final tag for the band. When DRAM address) all the tag data up to and including address endOfTagData has been read in, the te_finishedband signal is given and the doneBand flag is set. Read only. 0xDC FirstTagLineHeight 9 0 The number of lines minus 1 in the first tag encountered in this band. This will be equal to TagMaxLine if the band starts at a tag boundary. Read only. Work registers (set before starting the TE and must not be touched between bands) 0x100 LineInTag 1 0 Determines whether or not the first line of the page is in a line of tags or in an inter-tag gap. 1 - in a tag, 0 - in an inter-tag gap. 0x104 LinePos 14 0 The number of lines remaining minus 1, in either the tag or the inter-tag gap in at the start of the page. 0x110 to TagData 4 × 32 0 This 128 bit register must be 0x11C set up initially with the fixed data record for the page. This is either the lower 40 (or 56) bits (and the encodeFixed register should be set), or the lower 120 bits (and encodedFixed should be clear). The tagData[0] register contains the lower 32 bits and the tagData[3] register contains the upper 32 bits. This register is used throughout the tag encoding process to hold the next tag's variable data. Work registers (set internally) Read-only from the point of view of PCU register access 0x140 DotPos 14 0 Defines the number of dotpairs remaining in either the tag or inter-tag gap. Does not need to be setup. 0x144 CurrTagPlaneAdr 14 0 The dot-pair number being generated. 0x148 DotsInTag 1 0 Determines whether the current dot pair is in a tag or not 1 - in a tag, 0 - in an inter-tag gap. 0x14C TagAltSense 1 0 Determines whether the production of output dots is for the first (and subsequent even) or second (and subsequent odd) row of tags. 0x154 CurrTFSAdr 19 0 Points to the start next line of (64-bit aligned the TFS to be read in. DRAM address) 0x158 ReadsRemaining 4 0 Number of reads remaining in the current burst from the raw tag data interface 0x15C CountX 8 0 The number of tags remaining to be read (minus 1) by the raw tag data interface for the current line. 0x160 CountY 9 0 The number of times (minus 1) the tag data for the current line of tags needs to be read in by the raw tag data interface. 0x164 RtdTagSense 1 0 Determines whether the raw tag data interface is currently reading even rows of tags (=0) or odd rows of tags (=1) with respect to the start of the page. Note that this can be different from tagAltSense since the raw tag data interface is reading ahead of the production of dots. 0x168 RawTagDataAdr 19 0 The current read address (64-bit aligned within the unencoded raw DRAM address) tag data.

The PCU accessible registers are divided amongst the TE top level and the TE sub-blocks. This is achieved by including write decoders in the sub-blocks as well as the top level, see FIG. 189. In order to perform reads the sub-block registers are fed to the top level where the read decode is carried out on all the PCU accessible TE registers.

26.6.5.1 Starting the TE and Restarting the TE Between Bands

The TE must be started after the TFU.

For the first band of data, users set up NextBandStartTagDataAdr, NextBandEndTagData and NextBandFirstTagLineHeight as well as other TE configuration registers. Users then set the TE's Go bit to start processing of the band. When the tag data for the band has finished being decoded, the te_finishedband interrupt will be sent to the PCU and ICU indicating that the memory associated with the first band is now free. Processing can now start on the next band of tag data.

In order to process the next band NextBandStartTagDataAdr, NextBandEndTagData and NextBandFirstTagLineHeight need to be updated before writing a 1 to NextBandEnable. There are 4 mechanisms for restarting the TE between bands:

-   a. te_finishedband causes an interrupt to the CPU. The TE will have     set its DoneBand bit. The CPU reprograms the     NextBandStartTagDataAdr, NextBandEndTagData and     NextBandFirstTagLineHeight registers, and sets NextBandEnable to     restart the TE. -   b. The CPU programs the TE's NextBandStartTagDataAdr,     NextBandEndTagData and NextBandFirstTagLineHeight registers and sets     the NextBandEnable flag before the end of the current band. At the     end of the current band the TE sets DoneBand. As NextBandEnable is     already 1, the TE starts processing the next band immediately. -   c. The PCU is programmed so that te_finishedband triggers the PCU to     execute commands from DRAM to reprogram the NextBandStaifTagDataAdr,     NextBandEndTagData and NextBandFirstTagLineHeight registers and set     the NextBandEnable bit to start the TE processing the next band. The     advantage of this scheme is that the CPU could process band headers     in advance and store the band commands in DRAM ready for execution. -   d. This is a combination of b and c above. The PCU (rather than the     CPU in b) programs the TE's NextBandStartTagDataAdr,     NextBandEndTagData and NextBandFirstTagLineHeight registers and sets     the NextBandEnable bit before the end of the current band. At the     end of the current band the TE sets DoneBand and pulses     te_finishedband. As NextBandEnable is already 1, the TE starts     processing the next band immediately. Simultaneously,     te_finishedband triggers the PCU to fetch commands from DRAM. The TE     will have restarted by the time the PCU has fetched commands from     DRAM. The PCU commands program the TE next band shadow registers and     sets the NextBandEnable bit.

After the first tag on the page, all bands have their first tag start at the top i.e.

-   NextBandFirstTagLineHeight=TagMaxLine. Therefore the same value of     NextBandFirstTagLineHeight will normally be used for all bands.     Certainly, -   NextBandFirstTagLineHeight should not need to change after the     second time it is programmed.     26.6.6 TE Top Level FSM

The following diagram illustrates the states in the FSM.

At the highest level, the TE state machine steps through the output lines of a page one line at a time, with the starting position either in an inter-tag gap (signal dotsintag=0) or in a tag (signals tfsvalid and tdvalid and lineintag=1) (a SoPEC may be only printing part of a tag due to multiple SoPECs printing a single line).

If the current position is within an inter-tag gap, an output of 0 is generated. If the current position is within a tag, the tag format structure is used to determine the value of the output dot, using the appropriate encoded data bit from the fixed or variable data buffers as necessary. The TE then advances along the line of dots, moving through tags and inter-tag gaps according to the tag placement parameters.

Table 177 highlights the signals used within the FSM.

TABLE 177 Signals used within TE top level FSM Signal Name Function pclk Sync clock used to register all data within the FSM prst_n, te_reset Reset signals advtagline 1 cycles pulse indicating to TDI and TFS sub-blocks to move onto the next line of Tag data currdotlineadr[13:0] Address counter starting 2 pclk ahead of currtagplaneadr to generate the correct dotpair for the current line dotpos Counter to identify how many dotpairs wide the tag/gap is dotsintag Signal identifying whether the dotpair are in a tag(1)/gap(0) lineintag_temp Identical to lineintag but generated 1 pclk earlier linepos_shadow Shadow register for linepos due to linepos being written to by 2 different processes talaltsense Flag which alternates between tag/gap lines te_state FSM state variable teplanebuf 6-bit shift register used to format dotpairs into a byte for the TFU wradvline Advance line signal strobed when the last byte in a line is placed on te_tfu_wdata

Due to the 2 system clock delay in the TFS (both Table A and Table B outputs are registered) the TE FSM is working 2 system clock cycles AHEAD of the logic generating the write data for the TFU. As a result the following control signals had to be single/double registered on the system clock.

The tag_dot_line state can be broken down into 3 different stages.

Stage1:—The state tag_dot_line is entered due to the go signal becoming active. This state controls the writing of dotbytes to the TFU. As long as the tag line buffer address is not equal to the dotpairsperline register value and tfu_te_oktowrite is active, and there is valid TFS and TD available or taggaps, dotpairs are buffered into bytes and written to the TFU. The tag line buffer address is used internally but not supplied to the TFU since the TFU is a FIFO rather than the line store used in PEC1.

While generating the dotline of a tag/gap line (lineintag flag=1) the dot position counter dotpos is decremented/reloaded (with tagmaxdotpairs or taggapdot) as the TE moves between tags/gaps. The dotsintag flag is toggled between tags/gaps (0 for a gap, 1 for a tag). This pattern continues until the end of a dotline approaches (currdotlineadr==dotpairsperline).

2 system clock cycles before the end of the dotline the lineintag and tagaltsense signals must be prepared for the next dotline be it in a tag/gap dotline or a purely gap dotline.

Stage2:—At this point the end of a dot line is reached so it is time to decrement the linepos counter if still in a tag/gap row or reload the linepos register, dotpos counter and reprogram the dotsintag flag if going onto another tag/gap or pure gap row. Any signal with the _temp extension means this register is updated a cycle early in order for the real register to get its correct value while switching between dot lines and tag rows when dotpos and linepos counters reach zero i.e when dotpos=0 the end of a tag/gap has been reached, when linepos=0 the end of a tag row is reached. This stage uses the signals lineintag temp and tagaltsense which were generated one system clock cycle earlier in Stage 1.

Stage3:—This stage implements the writing of dotpairs to the correct part of the 6-bit shift register based on the LSBs of currtagplaneadr and also implements the counter for the currtagplaneadr. The currtagplaneadr is reset on reaching currtagplaneadr=(dotpairsperline−1). All the qualifier signals e.g dotsintag for this stage are delayed by 2 system clock cycles i.e. the currtagplaneadr (which is the internal write address not needed by the TFU) cannot be incremented until the dotpairs are available which is always 2 system clock cycles later than when currdotlineadr is incremented.

The wradvline and advtagline pulses are generated using the same logic (currently separated in the PEC1 Tag Encoder VHDL for clarity). Both of these pulses used to update further registers hence the reason they do not use the delayed by 2 system clock cycle qualifiers.

26.6.7 Combinational Logic

The TDI is responsible for providing the information data for a tag while the TFSI is responsible for deciding whether a particular dot on the tag should be printed as background pattern or tag information. Every dot within a tag's boundary is either an information dot or part of the background pattern.

The resulting lines of dots are stored in the TFU.

The TFSI reads one Tag Line Structure (TLS) from the DIU for every dot line of tags. Depending on the current printing position within the tag (indicated by the signal tagdotnum), the TFS interface outputs dot information for two dots and if necessary the corresponding read addresses for encoded tag data. The read address are supplied to the TDI which outputs the corresponding data values.

These data values (tdi_etd0 and tdi_etd1) are then combined with the dot information (tfsi_ta_dot0 and tfsi_ta_dot1) to produce the dot values that will actually be printed on the page (dots), see FIG. 192.

The signal lastdotintag is generated by checking that the dots are in a tag (dotsintag=1) and that the dotposition counter dotpos is equal to zero. It is also used by the TFS to load the index address register with zeros at the end of a tag as this is always the starting index when going from one tag to the next. lastdotintag is gated with advtagline in the TFSi (Table C) where adv_tfs_line pulse is used to update the Table C address reg for the new tag line—this is because lastdotintag occurs a cycle earlier than adv_tfs_line which would result in the wrong Table C value for the last dotpair. lastdotintag is also used in the TDi FSM (etd_switch state) to pulse the etd_advtag signal hence switching buffers in the ETDi for the next tag.

The signal lastdotintag1 is identical to lastdotintag except it is combinatorially generated (1 cycle earlier than lastdotintag, except at the end of a tagline). lastdotintag1 signal is only used in the TDi to reset the tdvalid signal on the cycle when dotpos=0. Note the UNSIGNED(currdotlineadr)=UNSIGNED(dotpairsperline)−1 not UNSIGNED(currdotlineadr)=UNSIGNED(dotpairsperline)−2 as in the lastdotintag_gen process as this is an combinatorial process.

The dotposvalid signal is created based on being in a tag line (lineintag1=1), dots being in a tag (dotsintag1=1), having a valid tag format structure available (tfsvalid1=1) and having encoded tag data available (tdvalid1=1). Note that each of the qualifier signals are delayed by 1 pclk cycle due to the registering of Table A output data into Table C where dotposvalid is used. The dotposvalid signal is used as an enable to load the Table C address register with the next index into Table B which in turn provides the 2 addresses to make 2 dots available.

The signal te_tfu_wdatavalid can only be active if in a taggap or if valid tag data is available (tdvalid2 and tfsvalid2) and the currtagpplaneadr(1:0) equal 11 i.e. a byte of data has been generated by combining four dotpairs.

The signal tagdotnum tells the TFS how many dotpairs remain in a tag/gap. It is calculated by subtracting the value in the dotpos counter from the value programmed in the tagmaxdotpairs register.

26.7 Tag Data Interface (TDI)

26.7.1 I/O Specification

TABLE 178 TDI Port List signal name I/O Description Clocks and Resets pclk In SoPEC system clock prst_n In Active-low, synchronous reset in pclk domain. DIU Read Interface Signals diu_data[63:0] In Data from DRAM. td_diu_rreq Out Data request to DRAM. td_diu_radr[21:5] Out Read address to DRAM. diu_td_rack In Data acknowledge from DRAM. diu_td_rvalid In Data valid signal from DRAM. PCU Interface Data, Control Signals and pcu_dataout[31:0] In PCU writes this data. pcu_addr[8:2] In PCU accesses this address. pcu_rwn In Global read/write-not signal from PCU. pcu_te_sel In PCU selects TE for r/w access. pcu_te_reset In PCU reset. td_te_doneband Out PCU readable registers. td_te_dataredun td_te_decode2den td_te_variabledatapresent td_te_encodefixed td_te_numtags0 td_te_numtags1 td_te_starttagdataadr td_te_rawtagdataadr td_te_endoftagdata td_te_firsttaglineheight td_te_tagdata0 td_te_tagdata1 td_te_tagdata2 td_te_tagdata3 td_te_countx td_te_county td_te_rtdtagsense td_te_readsremaining TFS (Tag Format Structure) tfsi_adr0[8:0] In Read address for dot0 tfsi_adr1[8:0] In Read address for dot1 Bandstore Signals cdu_startofbandstore[24:0] In Start memory area allocated for page bands cdu_endofbandstore[24:0] In Last address of the memory allocated for page bands te_finishedband Out Tag encoder band finished 26.7.2 Introduction

The tag data interface is responsible for obtaining the raw tag data and encoding it as required by the tag encoder. The smallest typical tag placement is 2 mm×2 mm, which means a tag is at least 126 1600 dpi dots wide.

In PEC1, in order to keep up with the HCU which processes 2 dots per cycle, the tag data interface has been designed to be capable of encoding a tag in 63 cycles. This is actually accomplished in approximately 52 cycles within PEC1. For SoPEC the TE need only produce one dot per cycle; it should be able to produce tags in no more than twice the time taken by the PEC1 TE. Moreover, any change in implementation from two dots to one dot per cycle should not lose the 63/52 cycle performance edge attained in the PEC1 TE.

As shown in FIG. 198, the tag data interface contains a raw tag data interface FSM that fetches tag data from DRAM, two symbol-at-a-time GF(2⁴) Reed-Solomon encoders, an encoded data interface and a state machine for controlling the encoding process. It also contains a tagData register that needs to be set up to hold the fixed tag data for the page.

The type of encoding used depends on the registers TE_encodefixed, TE_dataredun and TE_decode2den the options being,

-   -   (15,5) RS coding, where every 5 input symbols are used to         produce 15 output symbols, so the output is 3 times the size of         the input. This can be performed on fixed and variable tag data.     -   (15,7) RS coding, where every 7 input symbols are used to         produce 15 output symbols, so for the same number of input         symbols, the output is not as large as the (15,5) code (for more         details see section 26.7.6 on page 435). This can be performed         on fixed and variable tag data.     -   2D decoding, where each 2 input bits are used to produce 4         output bits. This can be performed on fixed and variable tag         data.     -   no coding, where the data is simply passed into the Encoded Data         Interface. This can be performed on fixed data only.

Each tag is made up of fixed tag data (i.e. this data is the same for each tag on the page) and variable tag data (i.e. different for each tag on the page).

Fixed tag data is either stored in DRAM as 120-bits when it is already coded (or no coding is required), 40-bits when (15,5) coding is required or 56-bits when (15,7) coding is required. Once the fixed tag data is coded it is 120-bits long. It is then stored in the Encoded Tag Data Interface. The variable tag data is stored in the DRAM in uncoded form. When (15,5) coding is required, the 120-bits stored in DRAM are encoded into 360-bits. When (15,7) coding is required, the 112-bits stored in DRAM are encoded into 240-bits. When 2D decoding is required the 120-bits stored in DRAM are converted into 240-bits. In each case the encoded bits are stored in the Encoded Tag Data Interface.

The encoded fixed and variable tag data are eventually used to print the tag.

The fixed tag data is loaded in once from the DRAM at the start of a page. It is encoded as necessary and is then stored in one of the 8×15-bits registers/RAMs in the Encoded Tag Data Interface. This data remains unchanged in the registers/RAMs until the next page is ready to be processed.

The 120-bits of unencoded variable tag data for each tag is stored in four 32-bit words. The TE re-reads the variable tag data, for a particular tag from DRAM, every time it produces that tag. The variable tag data FIFO which reads from DRAM has enough space to store 4 tags.

26.7.2.1 Bandstore Wrapping

Both TD and TFS storage in DRAM can wrap around the bandstore area. The bounds of the band store are described by inputs from the CDU shown in Table. The TD and TFS DRAM interfaces therefore support bandstore wrapping. If the TD or TFS DRAM interface increments an address it is checked to see if it matches the end of bandstore address. If so, then the address is mapped to the start of the bandstore.

26.7.3 Data Flow

An overview of the dataflow through the TDI can be seen in FIG. 198 below.

The TD interface consists of the following main sections:

-   -   the Raw Tag Data Interface—fetches tag data from DRAM;     -   the tag data register;     -   2 Reed Solomon encoders—each encodes one 4-bit symbol at a time;     -   the Encoded Tag Data Interface—supplies encoded tag data for         output;     -   Two 2D decoders.

The main performance specification for PEC1 is that the TE must be able to output data at a continuous rate of 2 dots per cycle.

26.7.4 Raw Tag Data Interface

The raw tag data interface (RTDI) provides a simple means of accessing raw tag data in DRAM. The RTDI passes tag data into a FIFO where it can be subsequently read as required. The 64-bit output from the FIFO can be read directly, with the value of the wr_rd_counter being used to set/reset as the enable signal (rtdAvail). The FIFO is clocked out with receipt of an rtdRd signal from the TS FSM.

FIG. 199 shows a block diagram of the raw tag data interface.

26.7.4.1 RTDI FSM

The RTDI state machine is responsible for keeping the raw tag FIFO full. The state machine reads the line of tag data once for each printline that uses the tag. This means a given line of tag data will be read TagHeight times. Typically this will be 126 times or more, based on an approximately 2 mm tag. Note that the first line of tag data may be read fewer times since the start of the page may be within a tag. In addition odd and even rows of tags may contain different numbers of tags. Section 26.6.5.1 outlines how to start the TE and restart it between bands. Users must set the NextBandStartTagDataAdr, NextBandEndOfTagData, NextBandFirstTagLineHeight and numTags[0], numTags[1] registers before starting the TE by asserting Go.

To restart the tag encoder for second and subsequent bands of a page, the NextBandStartTagDataAdr, NextBandEndOfTagData and NextBandFirstTagLineHeight registers need to be updated (typically numTags[0] and numTags[1] will be the same if the previous band contains an even number of tag rows) and NextBandEnable set. See Section 26.6.5.1 for a full description of the four ways of reprogramming the TE between bands.

The tag data is read once for every printline containing tags. When maximally packed, a row of tags contains 163 tags (see Table n page465 on page 408).

The RTDI State Flow diagram is shown in FIG. 200. An explanation of the states follows: idle state:—Stay in the idle state if there is no variable data present. If there is variable data present and there are at least 4 spaces left in the FIFO then request a burst of 2 tags from the DRAM (1*256 bits). Counter countx is assigned the number of tags in a even/odd line which depends on the value of register rtdtagsense. Down-counter county is assigned the number of dot lines high a tag will be (min 126). Initially it must be set the firsttaglineheight value as the TE may be between pages (i.e. a partial tag). For normal tag generation county will take the value of tagmaxline register.

diu_access:—The diu_access state will generate a request to the DRAM if there are at least 4 spaces in the FIFO. This is indicated by the counter wr_rd_counter which is incremented/decremented on writes/reads of the FIFO. As long as wr_rd_counter is less than 4 (FIFO is 8 high) there must be 4 locations free. A control signal called td_diu_radrvalid is generated for the duration of the DRAM burst access. Addresses are sent in bursts of 1. The counter burst_count controls this signal, (will involve modification to existing TE code.) If there is an odd number of tags in line then the last DRAM read will contain a tag in the first 128 bits and padding in the final 128 bits.

fifo_load:—This state controls the addressing to the DRAM. Counters countx and county are used to monitor whether the TE is processing a line of dots within a row of tags. When countx is zero it means all tag dots for this row are complete. When county is zero it means the TE is on the last line of dots (prior to Y scaling) for this row of tags. When a row of tags is complete the sense of rtdtagsense is inverted (odd/even). The rawtagdataadr is compared to the te_endoftagdata address. If rawtagdataadr=endoftagdata the doneband signal is set, the finishedband signal is pulsed, and the FSM enters the rtd_stall state until the doneband signal is reset to zero by the PCU by which time the rawtagdata, endoftagedata and firsttaglineheight registers are setup with new values to restart the TE. This state is used to count the 64-bit reads from the DIU. Each time diu_td_rvalid is high rtd_data_count is incremented by 1. The compare of rtd_data_count=rtd_num is necessary to find out when either all 4*64-bit data has been received or n*64-bit data (depending on a match of rawtagdataadr=endoftagdata in the middle of a set of 4*64-bit values being returned by the DIU.

rtd_stall:—This state waits for the the doneband signal to be reset (see page 426 for a description of how this occurs). Once reset the FSM returns to the idle state. This states also performs the same count on the diu_data read as above in the case where diu_td_rvalid has not gone high by the time the addressing is complete and the end of band data has been reached i.e. rawtagdataadr=endoftagdata

26.7.5 TDI State Machine

The tag data state machine has two processing phases. The first processing phase is to encode the fixed tag data stored in the 128-bit (2×64-bit) tag data register. The second is to encode tag data as it is required by the tag encoder.

When the Tag Encoder is started up, the fixed tag data is already preloaded in the 128 bit tag data record. If encodeFixed is set, then the 2 codewords stored in the lower bits of the tag data record need to be encoded: 40 bits if dataRedun=0, and 56 bits if dataRedun=1. If encodeFixed is clear, then the lower 120 bits of the tag data record must be passed to the encoded tag data interface without being encoded.

When encodeFixed is set, the symbols derived from codeword 0 are written to codeword 6 and the symbols derived from codeword 1 are written to codeword 7. The data symbols are stored first and then the remaining redundancy symbols are stored afterwards, for a total of 15 symbols. Thus, when dataRedun=0, the 5 symbols derived from bits 0-19 are written to symbols 0-4, and the redundancy symbols are written to symbols 5-14. When dataRedun=1, the 7 symbols derived from bits 0-27 are written to symbols 0-6, and the redundancy symbols are written to symbols 7-14.

When encodeFixed is clear, the 120 bits of fixed data is copied directly to codewords 6 and 7. The TDI State Flow diagram is shown in FIG. 202. An explanation of the states follows.

idle:—In the idle state wait for the tag encoder go signal—top_go=1. The first task is to either store or encode the Fixed data. Once the Fixed data is stored or encoded/stored the donefixed flag is set. If there is no variable data the FSM returns to the idle state hence the reason to check the donefixed flag before advancing i.e. only store/encode the fixed data once.

fixed_data:—In the fixed_data state the FSM must decode whether to directly store the fixed data in the ETDi or if the fixed data needs to be either (15:5) (40-bits) or (15:7) (56-bits) RS encoded or 2D decoded. The values stored in registers encodefixed and dataredun and decode2den determine what the next state should be.

bypass to etdi:—The bypass_to_etdi takes 120-bits of fixed data(pre-encoded) from the tag_data(127:0) register and stores it in the 15*8 (by 2 for simultaneous reads) buffers. The data is passed from the tag data register through 3 levels of muxing (level1, level2, level3) where it enters the RS0/RS1 encoders (which are now in a straight through mode (i.e. control_(—)5 and control_(—)7 are zero hence the data passes straight from the input to the output). The MSBs of the etd_wr_adr must be high to store this data as codewords 6,7.

etd_buf_switch:—This state is used to set the tdvalid signal and pulse the etd_adv_tag signal which in turn is used to switch the read write sense of the ETDi buffers (wrsb0). The firsttime signal is used to identify the first time a tag is encoded. If zero it means read the tag data from the RTDi FIFO and encode. Once encoded and stored the FSM returns to this state where it evaluates the sense of tdvalid. First time around it will be zero so this sets tdvalid and returns to the readtagdata state to fill the 2nd ETDi buffer. After this the FSM returns to this state and waits for the lastdotintag signal to arrive. In between tags when the lastdotingtag signal is received the etd_adv_tag is pulsed and the FSM goes to the readtagdata state. However if the lastdotintag signal arrives at the end of a line there is an extra 1 cycle delay introduced in generating the etd_adv_tag pulse (via etd_adv_tag_endofline) due to the pipelining in the TFS. This allows all the previous tag to be read from the correct buffer and seamless transfer to the other buffer for the next line.

readtagdata:—The readtagdata state waits to receive a rtdavail signal from the raw tag data interface which indicates there is raw tag data available. The tag_data register is 128-bits so it takes 2 pulses of the rtdrd signal to get the 2*64-bits into the tag_data register. If the rtdavail signal is set rtdrd is pulsed for 1 cycle and the FSM steps onto the loadtagdata state. Initially the flag first64 bits will be zero. The 64-bits of rtd are assigned to the tag_data[63:0] and the flag first64 bits is set to indicate the first raw tag data read is complete. The FSM then steps back to the read_tagdata state where it generates the second rtdrd pulse. The FSM then steps onto the loadtagdata state for where the second 64-bits of rawtag data are assigned to tag_data[128:64]. loadtagdata:—The loadtagdata state writes the raw tag data into the tag_data register from the RTDi FIFO. The first64 bits flag is reset to zero as the tag_data register now contains 120/112 bits of variable data. A decode of whether to (15:5) or (15:7) RS encode or 2D decode this data decides the next state.

rs_(—)15_(—)5:—The rs_(—)15_(—)5 (Reed Solomon (15:5) mode) state either encodes 40-bit Fixed data or 120-bit Variable data and provides the encoded tag data write address and write enable (etd_wr_adr and etdwe respectively). Once the fixed tag data is encoded the donefixed flag is set as this only needs to be done once per page. The variabledatapresent register is then polled to see if there is variable data in the tags. If there is variable data present then this data must be read from the RTDi and loaded into the tag_data register. Else the tdvalid flag must be set and FSM returns to the idle state. control_(—)5 is a control bit for the RS Encoder and controls feedforward and feedback muxes that enable (15:5) encoding.

The rs_(—)15_(—)5 state also generates the control signals for passing 120-bits of variable tag data to the RS encoder in 4-bit symbols per clock cycle. rs_counter is used both to control the level1_mux and act as the 15-cycle counter of the RS Encoder. This logic cycles for a total of 3*15 cycles to encode the 120-bits.

rs_(—)15_(—)7:—The rs_(—)15_(—)7 state is similar to the rs_(—)15_(—)5 state except the level1_mux has to select 7 4-bit symbols instead of 5.

decode_(—)2d 15_(—)5, decode_(—)2d_(—)15_(—)7:—The decode_(—)2d states provides the control signals for passing the 120-bit variable data to the 2D decoder. The 2 lsbs are decoded to create 4 bits. The 4 bits from each decoder are combined and stored in the ETDi. Next the 2 MSBs are decoded to create 4 bits. Again the 4 bits from each decoder are combined and stored in the ETDi.

As can be seen from Figure n page 488 on page Error! Bookmark not defined. there are 3 stages of muxing between the Tag Data register and the RS encoders or 2D decoders. Levels 1-2 are controlled by level1_mux and level2_mux which are generated within the TDi FSM as is the write address to the ETDi buffers (etd_wr_adr)

FIGS. 203 through 208 illustrate the mappings used to store the encoded fixed and variable tag data in the ETDI buffers.

26.7.6 Reed Solomon (RS) Encoder

26.7.7 Introduction

A Reed Solomon code is a non binary, block code. If a symbol consists of m bits then there are q=2^(m) possible symbols defining the code alphabet. In the TE, m=4 so the number of possible symbols is q=16.

An (n,k) RS code is a block code with k information symbols and n code-word symbols. RS codes have the property that the code word n is limited to at most q+1 symbols in length.

In the case of the TE, both (15,5) and (15,7) RS codes can be used. This means that up to 5 and 4 symbols respectively can be corrected.

Only one type of RS coder is used at any particular time. The RS coder to be used is determined by the registers TE_dataredun and TE_decode2den:

-   -   TE_dataredun=0 and TE_decode2den=0, then use the (15,5) RS coder     -   TE_dataredun=1 and TE_decode2den=0, then use the (15,7) RS coder

For a (15,k) RS code with m=4, k 4-bit information symbols applied to the coder produce 15 4-bit codeword symbols at the output. In the TE, the code is systematic so the first k codeword symbols are the same the as the k input information symbols.

A simple block diagram can be seen in.

26.7.8 I/O Specification

A I/O diagram of the RS encoder can be seen in.

26.7.9 Proposed Implementation

In the case of the TE, (15,5) and (15,7) codes are to be used with 4-bits per symbol.

The primitive polynomial is p(x)=x⁴x+x+1

In the case of the (15,5) code, this gives a generator polynomial of g(x)=(x+a)(x+a ²)(x+a ³)(x+a ⁴)(x+a ⁵)(x+a ⁶)(x+a ⁷)(x+a ⁸)(x+a ⁹)(x+a ¹⁰) g(x)=x ¹⁰ +a ² x ⁹ +a ³ x ⁸ +a ⁹ x ⁷ +a ⁶ x ⁶ +a ¹⁴ x ⁵ +a ² x ⁴ +ax ³ +a ⁶ x ² +ax+a ¹⁰ g(x)=x ¹⁰ +g ₉ x ⁹ +g ₈ x ⁸ +g ₇ x ⁷ +g ₆ x ⁶ +g ₅ x ⁵ +g ₄ x ⁴ +g ₃ x ³ +g ₂ x ² +g ₁ x+g ₀

In the case of the (15,7) code, this gives a generator polynomial of h(x)=(x+a)(x+a ²)(x+a ³)(x+a ⁴)(x+a ⁵)(x+a ⁶)(x+a ⁷)(x+a ⁸) h(x)=x ⁸ +a ¹⁴ x ⁷ +a ² x ⁶ +a ⁴ x ⁵ +a ² x ⁴ +a ¹³ x ³ +a ⁵ x ² +a ¹¹ x+a ⁶ h(x)=x ⁸ +h ₇ x ⁷ +h ₆ x ⁶ +h ₅ x ⁵ +h ₄ x ⁴ +h ₃ x ³ +h ₂ x ² +h ₁ x+h ₀

The output code words are produced by dividing the generator polynomial into a polynomial made up from the input symbols.

This division is accomplished using the circuit shown in FIG. 211.

The data in the circuit are Galois Field elements so addition and multiplication are performed using special circuitry. These are explained in the next sections.

The RS coder can operate either in (15,5) or (15,7) mode. The selection is made by the registers TE_dataredun and TE_decode2den.

When operating in (15,5) mode control_(—)7 is always zero and when operating in (15,7) mode control_(—)5 is always zero.

Firstly consider (15,5) mode i.e. TE_dataredun is set to zero.

For each new set of 5 input symbols, processing is as follows:

The 4-bits of the first symbol d₀ are fed to the input port rs_data_in(3:0) and control_(—)5 is set to 0. mux2 is set so as to use the output as feedback. control_(—)5 is zero so mux4 selects the input (rs_data_in) as the output (rs_data_out). Once the data has settled (<<1 cycle), the shift registers are clocked. The next symbol d₁ is then applied to the input, and again after the data has settled the shift registers are clocked again. This is repeated for the next 3 symbols d₂, d₃ and d₄. As a result, the first 5 outputs are the same as the inputs. After 5 cycles, the shift registers now contain the next 10 required outputs. control_(—)5 is set to 1 for the next 10 cycles so that zeros are fed back by mux2 and the shift register values are fed to the output by mux3 and mux4 by simply clocking the registers.

A timing diagram is shown below.

Secondly consider (15,7) mode i.e. TE_dataredun is set to one.

In this case processing is similar to above except that control_(—)7 stays low while 7 symbols (d₀, d₁ . . . d₆) are fed in. As well as being fed back into the circuit, these symbols are fed to the output.

After these 7 cycles, control_(—)7 is set to 1 and the contents of the shift registers are fed to the output.

A timing diagram is shown below.

The enable signal can be used to start/reset the counter and the shift registers.

The RS encoders can be designed so that encoding starts on a rising enable edge. After 15 symbols have been output, the encoder stops until a rising enable edge is detected. As a result there will be a delay between each codeword.

Alternatively, once the enable goes high the shift registers are reset and encoding will proceed until it is told to stop. rs_data_in must be supplied at the correct time. Using this method, data can be continuously output at a rate of 1 symbol per cycle, even over a few codewords.

Alternatively, the RS encoder can request data as it requires.

The performance criterion that must be met is that the following must be carried out within 63 cycles

-   -   load one tag's raw data into TE_tagdata     -   encode the raw tag data     -   store the encoded tag data in the Encoded Tag Data Interface

In the case of the raw fixed tag data at the start of a page, there is no definite performance criterion except that it should be encoded and stored as fast as possible.

26.7.10 Galois Field Elements and Their Representation

A Galois Field is a set of elements in which we can do addition, subtraction, multiplication and division without leaving the set.

The TE uses RS encoding over the Galois Field GF(2⁴). There are 2⁴ elements in GF(2⁴) and they are generated using the primitive polynomial p(x)=x⁴+x+1.

The 16 elements of GF(2⁴) can be represented in a number of different ways. Table 179 shows three possible representations—the power, polynomial and 4-tuple representation.

TABLE 179 GF(2⁴) representations 4-tuple power Polynomial representation representation Representation (a0 a1 a2 a3) 0 0 (0 0 0 0) 1 1 (1 0 0 0) A x (0 1 0 0) α²

x² (0 0 1 0) α³ x³ (0 0 0 1) α⁴ 1 + x (1 1 0 0) α⁵ x + x² (0 1 1 0) a⁶ x² + x³ (0 0 1 1) α⁷ 1 + x

+ x³ (1 1 0 1) α⁸ 1 + x² (1 0 1 0) α⁹ (0 1 0 1) x

+ x³ α¹⁰ 1 + x + x² (1 1 1 0) α¹¹ x + x² + x³ (0 1 1 1) α¹² 1 + x + x² + x³ (1 1 1 1) α¹³ 1 + x² + x³ (1 0 1 1) α¹⁴ 1 + x³ (1 0 0 1) 26.7.11 Multiplication of GF(2⁴) Elements

The multiplication of two field elements α^(a) and α^(b) is defined as

-   -   α¹=α^(a).α^(b)=α^((a+b)modulo 15)         Thus         α¹.α²=α³         α⁵.α¹⁰=α¹⁵         α⁶.α¹²=α³

So if we have the elements in exponential form, multiplication is simply a matter of modulo 15 addition.

If the elements are in polynomial/tuple form, the polynomials must be multiplied and reduced mod x⁴+x+1.

Suppose we wish to multiply the two field elements in GF(2⁴): α^(a) =a ₃ x ³ +a ₂ x ² +a ₁ x ¹ +a ₀ α^(b) =b ₃ x ³ +b ₂ x ² +b ₁ x ¹ +b ₀

-   -   where a₁, b_(i) are in the field (0,1) (i.e. modulo 2         arithmetic)

Multiplying these out and using x⁴+x+1=0 we get:

  α^(a + b) = [(a₀b₃ + a₁b₂ + a₂b₁ + a₃b₀) +   a₃b₃]x³ + [(a₀b₂ + a₁b₁ + a₂b₀) + a₃b₃ + (a₃b₂ + a₂b₃)]  x² + [(a₀b₁ + a₁b₀) + (a₃b₂ + a₂b₃) + (a₁b₃ + a₂b₂ +   a₃b₁)]x + [(a₀b₀ + a₁b₃ + a₂b₂ + a₃b₁)]    α^(a + b) = [a₀b₃ + a₁b₂ + a₂b₁ + a₃(b₀ + b₃)]x³ + [a₀b₂ + a₁b₁ +   a₂(b₀ + b₃) + a₃(b₂ + b₃)]x² + [a₀b₁ + a₁(b₀ + b₃) +   a₂(b₂ + b₃) + a₃(b₁ + b₂)]  x + [a₀b₀ + a₁b₃ + a₂b₂ + a₃b₁]

If we wish to multiply an arbitrary field element by a fixed field element we get a more simple form. Suppose we wish to multiply α^(b) by α³.

In this case α³=x³ so (a0 a1 a2 a3)=(0001). Substituting this into the above equation gives α^(c)=(b ₀ +b ₃)x ³+(b ₂ +b ₃)x ²+(b ₁ +b ₂)x+b ₁

This can be implemented using simple XOR gates as shown in FIG. 214

26.7.12 Addition of GF(2⁴) Elements

If the elements are in their polynomial/tuple form, polynomials are simply added.

Suppose we wish to add the two field elements in GF(2⁴): α⁸ =a ₃ x ³ +a ₂ x ² +a ₁ x+a ₀ α^(b) =b ₃ x ³ +b ₂ z ² +b ₁ x+b ₀

-   -   where a_(i), b_(i) are in the field (0,1) (i.e. modulo 2         arithmetic)         α^(c)=α^(a)+α^(b)=(a ₃ +b ₃)x ³+(a ₂ +b ₂)x2+(a ₁ +b ₁)x+(a ₀ +b         ₀)

Again this can be implemented using simple XOR gates as shown in FIG. 215

26.7.13 Reed Solomon Implementation

The designer can decide to create the relevant addition and multiplication circuits and instantiate them where necessary. Alternatively the feedback multiplications can be combined as follows. Consider the multiplication α^(a).α^(b)=α^(c)

-   -   or in terms of polynomials         (a ₃ x ³ +a ₂ x ² +a ₁ x+a ₀).(b ₃ x ³ +b ₂ x ² +b ₁ x+b ₀)=(c ₃         x ³ +c ₂ x ² +c ₁ x+c ₀)

If we substitute all of the possible field elements in for α^(a) and express α^(c) in terms of α^(b), we get the table of results shown in Table 180.

TABLE 180 α^(c) multiplied by all field elements, expressed in terms of α^(b) αa = a3x3 + a2x2 + a1x + a0 fixed field c3x3 + c2x2 + c1x + c0 element (a0 a1 a2 a3) c0 c1 c2 c3 0 (0 0 0 0) 1 (1 0 0 0) b₀ b₁ b₂ b₃ a (0 1 0 0) b₃ b₀ + b₃ b₁ b₂ α²  (0 0 1 0) b₂ b₂ + b₃ b₀ + b₃ b₁ α³  (0 0 0 1) b₁ b₁ + b₂ b₂ + b₃ b₀ + b₃ α⁴  (1 1 0 0) b₀ + b₃ b₀ + b₁ + b₃ b₁ + b₂ b₂ + b₃ α⁵  (0 1 1 0) b₂ + b₃ b₀ + b₂ b₀ + b₁ + b₃ b₁ + b₂ a⁶  (0 0 1 1) b₁ + b₂ b₁ + b₃ b₀ + b₂ b₀ + b₁ + b₃ α⁷  (1 1 0 1) b₀ + b₁ + b₃ b₀ + b₂ + b₃ b₁ + b₃ b₀ + b₂ α⁸  (1 0 1 0) b₀ + b₂ b₁ + b₂ + b₃ b₀ + b₂ + b₃ b₁ + b₃ α⁹  (0 1 0 1) b₁ + b₃ b₀ + b₁ + b₂ + b₃ b₁ + b₂ + b₃ b₀ + b₂ + b₃ α¹⁰ (1 1 1 0) b₀ + b₂ + b₃ b₀ + b₁ + b₂ b₀ + b₁ + b₂ + b₃ b₁ + b₂ + b₃ α¹¹ (0 1 1 1) b₁ + b₂ + b₃ b₀ + b₁ b₀ + b₁ + b₂ b₀ + b₁ + b₂ + b₃ α¹² (1 1 1 1) b₀ + b₁ + b₂ + b₃ b₀ b₀ + b₁ b₀ + b₁ + b₂ α¹³ (1 0 1 1) b₀ + b₁ + b₂ b₃ b₀ b₀ + b₁ α¹⁴ (1 0 0 1) b₀ + b₁ b₂ b₃ b₀ the following signals are required:

-   -   b₀, b₁, b₂, b₃,     -   (b₀+b₁), (b₀+b₂), (b₀+b₃), (b₁+b₂), (b₁+b₃), (b₂+b₃),     -   (b₀+b₁+b₂), (b₀+b₁+b₃), (b₀+b₂+b₃), (b₁+b₂+b₃),     -   (b₀+b₁+b₂+b₃)

The implementation of the circuit can be seen in Figure. The main components are XOR gates, 4-bit shift registers and multiplexers.

The RS encoder has 4 input lines labelled 0, 1, 2 & 3 and 4 output lines labelled 0, 1, 2 & 3. This labelling corresponds to the subscripts of the polynomial/4-tuple representation. The mapping of 4-bit symbols from the TE tagdata register into the RS is as follows:

-   -   the LSB in the TE_tagdata is fed into line0     -   the next most significant LSB is fed into line1     -   the next most significant LSB is fed into line2     -   the MSB is fed into line3

The RS output mapping to the Encoded tag data interface is similiar. Two encoded symbols are stored in an 8-bit address. Within these 8 bits:

-   -   line0 is fed into the LSB (bit 0/4)     -   line1 is fed into the next most significant LSB (bit 1/5)     -   line2 is fed into the next most significant LSB (bit 2/6)     -   line3 is fed into the MSB (bit 3/7)         267.14 2D Decoder

The 2D decoder is selected when TE_decode2den=1. It operates on variable tag data only. its function is to convert 2-bits into 4-bits according to Table 181.

TABLE 181 Operation of 2D decoder input output 0 0 0 0 0 1 0 1 0 0 1 0 1 0 0 1 0 0 1 1 1 0 0 0 26.7.15 Encoded Tag Data Interface

The encoded tag data interface contains an encoded fixed tag data store interface and an encoded variable tag data store interface, as shown in FIG. 217.

The two reord units simply reorder the 9 input bits to map low-order codewords into the bit selection component of the address as shown in Table 182. Reordering of write addresses is not necessary since the addresses are already in the correct format.

TABLE 182 Reord unit input output bit# bit interpretation bit interpretation 8 A select 1 of 8 A select 1 of 4 codewords codeword tables 7 B B 6 C D select 1 of 15 symbols 5 D select 1 of 15 E symbols 4 E F 3 F G 2 G C select 1 of 8 bits 1 H select 1 of 4 H bits 0 I I

The encoded fixed data interface is a single 15×8-bit RAM with 2 read ports and 1 write port. As it is only written to during page setup time (it is fixed for the duration of a page) there is no need for simultaneous read/write access. However the fixed data store must be capable of decoding two simultaneous reads in a single cycle. FIG. 218 shows the implementation of the fixed data store.

The encoded variable tag data interface is a double buffered 3×15×8-bit RAM with 2 read ports and 1 write port. The double buffering allows one tag's data to be read (two reads in a single cycle) while the next tag's variable data is being stored. Write addressing is 6 bits: 2 bits of address for selecting 1 of 3, and 4 bits of address for selecting 1 of 15. Read addressing is the same with the addition of 3 more address bits for selecting 1 of 8.

FIG. 219 shows the implementation of the encoded variable tag data store. Double buffering is implemented via two sub-buffers. Each time an AdvTag pulse is received, the sense of which sub-buffer is being read from or written to changes. This is accomplished by a 1-bit flag called wrsb0. Although the initial state of wrsb0 is irrelevant, it must invert upon receipt of an AdvTag pulse. The structure of each sub-buffer is shown in FIG. 220.

26.8 Tag Format Structure (TFS) Interface

26.8.1 Introduction

The TFS specifies the contents of every dot position within a tags border i.e.:

-   -   is the dot part of the background?     -   is the dot part of the data?

The TFS is broken up into Tag Line Structures (TLS) which specify the contents of every dot position in a particular line of a tag. Each TLS consists of three tables—A, B and C (see FIG. 221).

For a given line of dots, all the tags on that line correspond to the same tag line structure.

Consequently, for a given line of output dots, a single tag line structure is required, and not the entire TFS. Double buffering allows the next tag line structure to be fetched from the TFS in DRAM while the existing tag line structure is used to render the current tag line.

The TFS interface is responsible for loading the appropriate line of the tag format structure as the tag encoder advances through the page. It is also responsible for producing table A and table B outputs for two consecutive dot positions in the current tag line.

-   -   There is a TLS for every dot line of a tag.     -   All tags that are on the same line have the exact same TLS.     -   A tag can be up to 384 dots wide, so each of these 384 dots must         be specified in the TLS.     -   The TLS information is stored in DRAM and one TLS must be read         into the TFS Interface for each line of dots that are outputted         to the Tag Plane Line Buffers.     -   Each TLS is read from DRAM as 5 times 256-bit words with 214         padded bits in the last 256-bit DRAM read.         26.8.2 I/O Specification

TABLE 183 Tag Format Structure Interface Port List signal signal name type description Pclk In SoPEC system clock prst_n In Active-low, synchronous reset in pclk domain top_go In Go signal from TE top level DRAM diu_data[63:0] In Data from DRAM diu_tfs_rack In Data acknowledge from DRAM diu_tfs_rvalid In Data valid from DRAM tfs_diu_rreq Out Read request to DRAM tfs_diu_radr[21:5] Out Read address to DRAM tag encoder top level top_advtagline In Pulsed after the last line of a row of tags top_tagaltsense In For even tag rows = 0 i.e. 0, 2, 4 . . . For odd tag rows = 1 i.e. 1, 3, 5 . . . top_lastdotintag In Last dot in tag is currently being processed top_dotposvalid In Current dot position is a tag dot and its structure data and tag data is available top_tagdotnum[7:0] In Counts from zero up to TE_tagmaxdotpairs (min. = 1, max. = 192) tfsi_valid Out TLS tables A, B and C, ready for use tfsi_ta_dot0[1:0] Out Even entry from Table A corresponding to top_tagdotnum tfsi_ta_dot1[1:0] Out Odd entry from Table A corresponding to top_tagdotnum tag encoder top level (PCU read decoder) tfs_te_tfsstartadr[23:0] Out TFS tfsstartadr register tfs_te_tfsendadr[23:0] Out TFS tfsendadr register tfs_te_tfsfirstlineadr[23:0] Out TFS tfsfirstlineadr register tfs_te_currtfsadr[23:0] Out TFS currtfsadr register TDI tfsi_tdi_adr0[8:0] Out Read address for dot0 (even dot) tfsi_tdi_adr1[8:0] Out Read address for dot1 (odd dot) 26.8.2.1 State Machine

The state machine is responsible for generating control signals for the various TFS table units, and to load the appropriate line from the TFS. The states are explained below.

idle:—Wait for top_go to become active. Pulse adv_tfs_line for 1 cycle to reset tawradr and tbwradr registers. Pulsing adv_tfs_line will switch the read/write sense of Table B so switching Table A here as well to keep things the same i.e. wrta0=NOT(wrta0).

diu_access:—In the diu_access state a request is sent to the DIU. Once an ack signal is received Table A write enable is asserted and the FSM moves to the tls_load state.

tls_load:—The DRAM access is a burst of 5 256-bit accesses, ultimately returned by the DIU as 5*(4*64 bit) words. There will be 192 padded bits in the last 256-bit DRAM word. The first 12 64-bit words reads are for Table A, words 12 to 15 and some of 16 are for Table B while part of read 16 data is for Table C. The counter read_num is used to identify which data goes to which table. The table B data is stored temporarily in a 288-bit register until the tls_update state hence tbwe does not become active until read_num=16).

-   -   The DIU data goes directly into Table A (12*64).     -   The DIU data for Table B is loaded into a 288-bit register.     -   The DIU data goes directly into Table C.

tls_update:—The 288-bits in Table B need to written to a 32*9 buffer. The tls_update state takes care of this using the read_num counter.

tls_next—This state checks the logic level of tfsvalid and switches the read/write senses of Table A (wrta0) and Table B a cycle later (using the adv_tfs_line pulse). The reason for switching Table A a cycle early is to make sure the top_level address via tagdotnum is pointing to the correct buffer. Keep in mind the top_level is working a cycle ahead of Table A and 2 cycles ahead of Table B.

If tfsValid is 1, the state machine waits until the advTagLine signal is received. When it is received, the state machine pulses advTFSLine (to switch read/write sense in tables A, B, C), and starts reading the next line of the TFS from currTFSAdr.

If tfsValid is 0, the state machine pulses advTFSLine (to switch read/write sense in tables A, B, C) and then jumps to the tls_tfsvalid_set state where the signal tfsValid is set to 1 (allowing the tag encoder to start, or to continue if it had been stalled). The state machine can then start reading the next line of the TFS from currTFSAdr.

tls_tfsvalid_next:—Simply sets the tfsvalid signal and returns the FSM to the diu_access state.

If an advTagLine signal is received before the next line of the TFS has been read in, tfsValid is cleared to 0 and processing continues as outlined above.

26.8.2.2 Bandstore Wrapping

Both TD and TFS storage in DRAM can wrap around the bandstore area. The bounds of the band store are described by inputs from the CDU shown in Table. The TD and TFS DRAM interfaces therefore support bandstore wrapping. If the TD or TFS DRAM interface increments an address it is checked to see if it matches the end of bandstore address. If so, then the address is mapped to the start of the bandstore.

The TFS state flow diagram is shown in below.

26.8.3 Generating a Tag From Tables A, B and C

The TFS contains an entry for each dot position within the tag's bounding box. Each entry specifies whether the dot is part of the constant background pattern or part of the tag's data component (both fixed and variable).

The TFS therefore has TagHeight×TagWidth entries, where TagHeight is the height of the tag in dot-lines and TagWidth is the width of the tag in dots. The TFS entries that specify a single dot-line of a tag are known as a Tag Line Structure.

The TFS contains a TLS for each of the 1600 dpi lines in the tag's bounding box. Each TLS contains three contiguous tables, known as tables A, B and C.

Table A contains 384 2-bit entries i.e. one entry for each dot in a single line of a tag up to the maximum width of a tag. The actual number of entries used should match the size of the bounding box for the tag in the dot dimension, but all 384 entries must be present.

Table B contains 32 9-bit data address that refer to (in order of appearance) the data dots present in the particular line. Again, all 32 entries must be present, even if fewer are used.

Table C contains two 5-bit pointers into table B and is followed by 22 unused bits. The total length of each TLS is therefore 34 32-bit words.

Each output dot value is generated as follows: Each entry in Table A consists of 2-bits-bit0 and bit1. These 2-bits are interpreted according to Table 184, Table 185 and Table 186.

TABLE 184 Interpretation of bit0 from entry in Table A bit0 interpretation 0 the output bit comes directly from bit1 (see Table ). 1 the output bit comes from a data bit. Bit1 is used in conjunction with Tag Line Structure Table B to determine which data bit will be output.

TABLE 185 Interpretation of bit1 from entry in table A when bit0 = 0 bit 1 interpretation 0 output 0 1 output 1

TABLE 186 Interpretation of bit1 from entry in table A when bit0 = 1 bit 1 interpretation 0 output data bit pointed to by current index into Table B. 1 output data bit pointed to by current index into Table B, and advance index by 1.

If bit0=0 then the output dot for this entry is part of the constant background pattern. The dot value itself comes from bit1 i.e. if bit1=0 then the output is 0 and if bit1=1 then the output is 1. If bit0=1 then the output dot for this entry comes from the variable or fixed tag data. Bit1 is used in conjunction with Tables B and C to determine data bits to use.

To understand the interpretation of bit1 when bit0=1 we need to know what is stored in Table B. Table B contains the addresses of all the data bits that are used in the particular line of a tag in order of appearance. Therefore, up to 32 different data bits can appear in a line of a tag. The address of the first data dot in a tag will be given by the address stored in entry 0 of Table B. As we advance along the various data dots we will advance through the various Table B entries. Each Table B entry is 9-bits long and each points to a specific variable or fixed data bit for the tag. Each tag contains a maximum of 120 fixed and 360 variable data bits, for a total of 480 data bits. To aid address decoding, the addresses are based on the RS encoded tag data. Table lists the interpretation of the 9-bit addresses.

TABLE 187 Interpretation of 9-bit tag data address in Table B bit pos name description 8 CodeWordSelect Select 1 of 8 codewords. Codewords 0, 1, 2, 3, 4, 5 are variable data. Codewords 6, 7 are fixed data. 7 6 5 SymbolSelect Select 1 of 15 symbols (1111 invalid) 4 3 2 1 BitSelect Select 1 of 4 bits from the selected symbols 0

If the fixed data is supplied to the TE in an unencoded form, the symbols derived from codeword 0 of fixed data are written to codeword 6 and the symbols derived from fixed data codeword 1 are written to codeword 7. The data symbols are stored first and then the remaining redundancy symbols are stored afterwards, for a total of 15 symbols. Thus, when 5 data symbols are used, the 5 symbols derived from bits 0-19 are written to symbols 0-4, and the redundancy symbols are written to symbols 5-14. When 7 data symbols are used, the 7 symbols derived from bits 0-27 are written to symbols 0-6, and the redundancy symbols are written to symbols 7-14

However, if the fixed data is supplied to the TE in a pre-encoded form, the encoding could theoretically be anything. Consequently the 120 bits of fixed data is copied to codewords 6 and 7 as shown in Table 188.

TABLE 188 Mapping of fixed data to codeword/symbols when no redundancy encoding output output input bits symbol range codeword  0-19 0-4 6 20-39 0-4 7 40-59 5-9 6 60-79 5-9 7 80-99 10-14 6 100-119 10-14 7

It is important to note that the interpretation of bit1 from Table A (when bit0-1) is relative. A 5-bit index is used to cycle through the data address in Table B. Since the first tag on a particular line may or may not start at the first dot in the tag, an initial value for the index into Table B is needed. Subsequent tags on the same line will always start with an index of 0, and any partial tag at the end of a line will simply finish before the entire tag has been rendered. The initial index required due to the rendering of a partial tag at the start of a line is supplied by Table C. The initial index will be different for each TLS and there are two possible initial indexes since there are effectively two types of rows of tags in terms of initial offsets.

Table C provides the appropriate start index into Table B (2 5-bit indices). When rendering even rows of tags, entry 0 is used as the initial index into Table B, and when rendering odd rows of tags, entry 1 is used as the initial index into Table B. The second and subsequent tags start at the left most dots position within the tag, so can use an initial index of 0.

26.8.4 Architecture

A block diagram of the Tag Format Structure Interface can be seen in FIG. 223.

26.8.4.1 Table A Interface

The implementation of table A is two 16×64-bit RAMs with a small amount of control logic, as shown in FIG. 224. While one RAM is read from for the current line's table A data (4 bits representing 2 contiguous table A entries), the other RAM is being written to with the next line's table A data (64-bits at a time).

Note:—The Table A data to be printed (if each LSB=0) must be passed to the top_level 2 cycles after the read of Table A due to the 2-stage pipelining in the TFS from registering Table A and Table B outputs hence this extra registering stage for the generation of ta_dot0_(—)1 cyclelater and ta_dot1_(—)1 cyclelater.

Each time an AdvTFSLine pulse is received, the sense of which RAM is being read from or written to changes. This is accomplished by a 1-bit flag called wrta0. Although the initial state of wrta0 is irrelevant, it must invert upon receipt of an AdvTFSLine pulse. A 4-bit counter called taWrAdr keeps the write address for the 12 writes that occur after the start of each line (specified by the AdvTFSLine control input). The tawe (table A write enable) input is set whenever the data in is to be written to table A. The taWrAdr address counter automatically increments with each write to table A. Address generation for tawe and taWrAdr is shown in Table 189.

26.8.4.2 Table C Interface

A block diagram of the table C interface is shown below in FIG. 226.

The address generator for table C contains a 5 bit address register adr that is set to a new address at the start of processing the tag (either of the two table C initial values based on tagAltSense at the start of the line, and 0 for subsequent tags on the same line). Each cycle two addresses into table B are generated based on the two 2-bit inputs (in0 and in1). As shown in Section 189, the output address tbRdAdr0 is always adr and tbRdAdr1 is one of adr and adr+1, and at the end of the cycle adr takes on one of adr, adr+1, and adr+2.

TABLE 189 AdrGen lookup table inputs outputs in0 in1 adr0Sel adr1Sel adrSel 00 00 X¹⁸ X adr 00 01 X adr adr 00 10 X X adr 00 11 X adr adr + 1 01 00 adr X adr 01 01 adr adr adr 01 10 adr X adr 01 11 adr adr adr + 1 10 00 X X adr 10 01 X adr adr 10 10 X X adr 10 11 X adr adr + 1 11 00 adr X adr + 1 11 01 adr adr + 1 adr + 1 11 10 adr X adr + 1 11 11 adr adr + 1 adr + 2 ¹⁸X = don't care state 26.8.4.3 Table B Interface

The table B interface implementation generates two encoded tag data addresses (tfsi_adr0, tfsu_adr1) based on two table B input addresses (tbRdAdr0, tbRdAdr1). A block diagram of table B can be seen in FIG. 227.

Table B data is initially loaded into the 288-bit table B temporary register via the TFS FSM. Once all 288-bit entries have been loaded from DRAM, the data is written in 9-bit chunks to the 32*9 register arrays based on tbwradr.

Each time an AdvTFSLine pulse is received, the sense of which sub buffer is being read from or written to changes. This is accomplished by a 1-bit flag called wrtb0. Although the initial state of wrtb0 is irrelevant, it must invert upon receipt of an AdvTFSLine pulse.

Note:—The output addresses from Table B are registered.

27 Tag FIFO Unit (TFU)

27.1 Overview

The Tag FIFO Unit (TFU) provides the means by which data is transferred between the Tag Encoder (TE) and the HCU. By abstracting the buffering mechanism and controls from both units, the interface is clean between the data user and the data generator.

The TFU is a simple FIFO interface to the HCU. The Tag Encoder will provide support for arbitrary Y integer scaling up to 1600 dpi. X integer scaling of the tag dot data is performed at the output of the FIFO in the TFU. There is feedback to the TE from the TFU to allow stalling of the TE during a line. The TE interfaces to the TFU with a data width of 8 bits. The TFU interfaces to the HCU with a data width of 1 bit.

The depth of the TFU FIFO is chosen as 16 bytes so that the FIFO can store a single 126 dot tag.

27.1.1 Interfaces Between TE, TFU and HCU

27.1.1.1 TE-TFU Interface

The interface from the TE to the TFU comprises the following signals:

-   -   te_tfu_wdata, 8-bit write data.     -   te_tfu_wdatavalID, write data valid.     -   te_tfu_wradvline, accompanies the last valid 8-bit write data in         a line.

The interface from the TFU to TE comprises the following signal:

-   -   tfu_te_oktowrite, indicating to the TE that there is space         available in the TFU FIFO.

The TE writes data to the TFU FIFO as long as the TFU's tfu_te_oktowrite output bit is set. The TE write will not occur unless data is accompanied by a data valid signal.

27.1.1.2 TFU-HCU Interface

The interface from the TFU to the HCU comprises the following signals:

-   -   tfu_hcu_tdata, 1-bit data.     -   tfu_hcu_avail, data valid signal indicating that there is data         available in the TFU FIFO.

The interface from HCU to TFU comprises the following signal:

-   -   hcu_tfu_ready, indicating to the TFU to supply the next dot.         27.1.1.2.1 X Scaling

Tag data is replicated a scale factor (SF) number of times in the X direction to convert the final output to 1600 dpi. Unlike both the CFU and SFU, which support non-integer scaling, the scaling is integer only. Replication in the X direction is performed at the output of the TFU FIFO on a dot-by-dot basis.

To account for the case where there may be two SoPEC devices, each generating its own portion of a dot-line, the first dot in a line may not be replicated the total scale-factor number of times by an individual TFU. The dot will ultimately be scaled-up correctly with both devices doing part of the scaling, one on its lead-out and the other on its lead in.

Note two SoPEC TEs may be involved in producing the same byte of output tag data straddling the printhead boundary. The HCU of the left SoPEC will accept from its TE the correct amount of dots, ignoring any dots in the last byte that do not apply to its printhead. The TE of the right SoPEC will be programmed the correct number of dots into the tag and its output will be byte aligned with the left edge of the printhead.

27.2 Definitions OF I/O

TABLE 190 TFU Port List Port Name Pins I/O Description Clocks and Resets Pclk 1 In SoPEC Functional clock. Prst_n 1 In Global reset signal. PCU Interface data and control signals Pcu_adr[4:2] 2 In PCU address bus. Only 3 bits are required to decode the address space for this block. Pcu_dataout[31:0] 32 In Shared write data bus from the PCU. Tfu_pcu_datain[31:0] 32 Out Read data bus from the TFU to the PCU. Pcu_rwn 1 In Common read/not-write signal from the PCU. Pcu_tfu_sel 1 In Block select from the PCU. When pcu_tfu_sel is high both pcu_adr and pcu_dataout are valid. Tfu_pcu_rdy 1 Out Ready signal to the PCU. When tfu_pcu_rdy is high it indicates the last cycle of the access. For a write cycle this means pcu_dataout has been registered by the block and for a read cycle this means the data on tfu_pcu_datain is valid. TE Interface data and control signals Te_tfu_wdata[7:0] 8 In Write data for TFU FIFO. Te_tfu_wdatavalid 1 In Write data valid signal. Te_tfu_wradvline 1 In Advance line signal strobed when the last byte in a line is placed on te_tfu_wdata tfu_te_oktowrite 1 Out Ready signal indicating TFU has space available in it's FIFO and is ready to be written to. HCU Interface data and control signals Hcu_tfu_advdot 1 In Signal indicating to the TFU that the HCU is ready to accept the next dot of data from TFU. tfu_hcu_tdata 1 Out Data from the TFU FIFO. tfu_hcu_avail 1 Out Signal indicating valid data available from TFU FIFO. 27.3 Configuration Registers

TABLE 191 TFU Configuration Registers value Address register # on TFU_Base+ name bits reset description Control registers 0x00 Reset 1 1 A write to this register causes a reset of the TFU. This register can be read to indicate the reset state: 0 - reset in progress 1 - reset not in progress. 0x04 Go 1 see Writing 1 to this register text starts the TFU. Writing 0 to this register halts the TFU. When Go is deasserted the state-machines go to their idle states but all count- ers and configuration registers keep their values. When Go is asserted all counters are reset, but configuration registers keeptheir values (i.e. they don't get reset). The TFU must be started before the TE is started. This register can be read to determine if the TFU is running (1 = running, 0 = stopped). Setup registers (constant during processing of page) 0x08 XScale 8 1 Tag scale factor in X direction. 0x0C XFracScale 8 1 Tag scale factor in X direction for the first dot in a line (must be programmed to be less than or equal to XScale) 0x10 TEByteCount 12 0 The number of bytes to be accepted from the TE per line. Once this number of bytes have been re- ceived subsequent bytes are ignored until there is a strobe on the te_tfu_wradvline 0x14 HCUDotCount 16 0 The number of (optionally) x-scaled dots per line to be supplied to the HCU. Once this number has been reached the remainder of the current FIFO byte is ignored. 27.4 Detailed Description

The FIFO is a simple 16-byte store with read and write pointers, and a contents store, FIG. 229. 16 bytes is sufficient to store a single 126 dot tag.

Each line a total of TEByteCount bytes is read into the FIFO. All subsequent bytes are ignored until there is a strobe on the te_tfu_wradvline signal, whereupon bytes for the next line are stored. On the HCU side, a total of HCUDotCount dots are produced at the output. Once this count is reached any more dots in the FIFO byte currently being processed are ignored. For the first dot in the next line the start of line scale factor, XFracScale, is used.

The behaviour of these signals and the control signals between the TFU and the TE and HCU is detailed below.

// Concurrently Executed Code: // TE always allowed to write when there's either (a) room  or  (b)  no  room  and  all // bytes for that line have been received. if ((FifoCntnts != FifoMax) OR (FifoCntnts == FifoMax and ByteToRx == 0)) then tfu_te_oktowrite = 1 else tfu_te_oktowrite = 0 // Data presented to HCU when there is (a) data in FIFO  and  (b)  the  HCU  has  not // received all dots for a line if (FifoCntnts != 0) AND (BitToTx != 0)then tfu_hcu_avail = 1 else tfu_hcu_avail = 0 // Output mux of FIFO data tfu_hcu_tdata = Fifo[FifoRdPnt] [RdBit] // Sequentially Executed Code: if  (te_tfu_wdatavalid  ==  1)  AND  (FifoCntnts  != FifoMax) AND (ByteToRx != 0) then Fifo[FifoWrPnt] = te_tfu_wdata FifoWrPnt ++ FifoContents ++ ByteToRx −− if (te_tfu_wradvline == 1) then ByteToRx = TEByteCount if (hcu_tfu_advdot == 1 and FifoCntnts != 0) then { BitToTx ++ if (RepFrac == 1) then RepFrac = Xscale if (RdBit = 7) then RdBit = 0 FifoRdPnt ++ FifoContents −− else RdBit++ else RepFrac−− if (BitToTx == 1) then { RepFrac = XFracScale RdBit = 0 FifoRdPnt ++ FifoContents−− BitToTx = HCUDotCount } }

What is not detailed above is the fact that, since this is a circular buffer, both the fifo read and write-pointers wrap-around to zero after they reach two. Also not detailed is the fact that if there is a change of both the read and write-pointer in the same cycle, the fifo contents counter remains unchanged.

28 Alftoner Compositor Unit (HCU)

28.1 Overview

The Halftoner Compositor Unit (HCU) produces dots for each nozzle in the destination printhead taking account of the page dimensions (including margins). The spot data and tag data are received in bi-level form while the pixel contone data received from the CFU must be dithered to a bi-level representation. The resultant 6 bi-level planes for each dot position on the page are then remapped to 6 output planes and output dot at a time (6 bits) to the next stage in the printing pipeline, namely the dead nozzle compensator (DNC).

28.2 Data Flow

FIG. 230 shows a simple dot data flow high level block diagram of the HCU. The HCU reads contone data from the CFU, bi-level spot data from the SFU, and bi-level tag data from the TFU.

Dither matrices are read from the DRAM via the DIU. The calculated output dot (6 bits) is read by the DNC.

The HCU is given the page dimensions (including margins), and is only started once for the page. It does not need to be programmed in between bands or restarted for each band. The HCU will stall appropriately if its input buffers are starved. At the end of the page the HCU will continue to produce 0 for all dots as long as data is requested by the units further down the pipeline (this allows later units to conveniently flush pipelined data).

The HCU performs a linear processing of dots calculating the 6-bit output of a dot in each cycle. The mapping of 6 calculated bits to 6 output bits for each dot allows for such example mappings. as compositing of the spot0 layer over the appropriate contone layer (typically black), the merging of CMY into K (if K is present in the printhead), the splitting of K into CMY dots if there is no K in the printhead, and the generation of a fixative output bitstream.

28.3 Dram Storage Requirements

SoPEC allows for a number of different dither matrix configurations up to 256 bytes wide. The dither matrix is stored in DRAM. Using either a single or double-buffer scheme a line of the dither matrix must be read in by the HCU over a SoPEC line time. SoPEC must produce 13824 dots per line for A4/Letter printing which takes 13824 cycles.

The following give the storage and bandwidths requirements for some of the possible configurations of the dither matrix.

-   -   4 Kbyte DRAM storage required for one 64×64 (preferred) byte         dither matrix     -   6.25 Kbyte DRAM storage required for one 80×80 byte dither         matrix     -   16 Kbyte DRAM storage required for four 64×64 byte dither         matrices     -   64 Kbyte DRAM storage required for one 256×256 byte dither         matrix

It takes 4 or 8 read accesses to load a line of dither matrix into the dither matrix buffer, depending on whether we're using a single or double buffer (configured by DoubleLineBuff register).

28.4 Implementation

A block diagram of the HCU is given in FIG. 231.

28.4.1 Definition of I/O

TABLE 192 HCU port list and description Port name Pins I/O Description Clocks and reset Pclk 1 In System clock. prst_n 1 In System reset, synchronous active low. PCU interface pcu_hcu_sel 1 In Block select from the PCU. When pcu_hcu_sel is high both pcu_adr and pcu_dataout are valid. pcu_rwn 1 In Common read/not-write signal from the PCU. pcu_adr[7:2] 6 In PCU address bus. Only 6 bits are required to decode the address space for this block. pcu_dataout[31:0] 32 In Shared write data bus from the PCU. hcu_pcu_rdy 1 Out Ready signal to the PCU. When hcu_pcu_rdy is high it indicates the last cycle of the access. For a write cycle this means pcu_dataout has been registered by the block and for a read cycle this means the data on hcu_pcu_datain is valid. hcu_pcu_datain[31:0] 32 Out Read data bus to the PCU. DIU interface hcu_diu_rreq 1 Out HCU read request, active high. A read request must be accompanied by a valid read address. diu_hcu_rack 1 In Acknowledge from DIU, active high. Indicates that a read request has been accepted and the new read address can be placed on the address bus, hcu_diu_radr. hcu_diu_radr[21:5] 17 Out HCU read address. 17 bits wide (256-bit aligned word). diu_hcu_rvalid 1 In Read data valid, active high. Indicates that valid read data is now on the read data bus, diu_data. diu_data[63:0] 64 In Read data from DIU. CFU interface cfu_hcu_avail 1 In Indicates valid data present on cfu_hcu_c[3-0]data lines. cfu_hcu_c0data[7:0] 8 In Pixel of data in contone plane 0. cfu_hcu_c1data[7:0] 8 In Pixel of data in contone plane 1. cfu_hcu_c2data[7:0] 8 In Pixel of data in contone plane 2. cfu_hcu_c3data[7:0] 8 In Pixel of data in contone plane 3. hcu_cfu_advdot 1 Out Informs the CFU that the HCU has captured the pixel data on cfu_hcu_c[3-0]data lines and the CFU can now place the next pixel on the data lines. SFU interface sfu_hcu_avail 1 In Indicates valid data present on sfu_hcu_sdata. sfu_hcu_sdata 1 In Bi-level dot data. hcu_sfu_advdot 1 Out Informs the SFU that the HCU has captured the dot data on sfu_hcu_sdata and the SFU can now place the next dot on the data line. TFU interface tfu_hcu_avail 1 In Indicates valid data present on tfu_hcu_tdata. tfu_hcu_tdata 1 In Tag dot data. hcu_tfu_advdot 1 Out Informs the TFU that the HCU has captured the dot data on tfu_hcu_tdata and the TFU can now place the next dot on the data line. DNC interface dnc_hcu_ready 1 In Indicates that DNC is ready to accept data from the HCU. hcu_dnc_avail 1 Out Indicates valid data present on hcu_dnc_data. hcu_dnc_data[5:0] 6 Out Output bi-level dot data in 6 ink planes 28.4.2 Configuration Registers

The configuration registers in the HCU are programmed via the PCU interface. Refer to section 21.8.2 on page 321 for the description of the protocol and timing diagrams for reading and writing registers in the HCU. Note that since addresses in SoPEC are byte aligned and the PCU only supports 32-bit register reads and writes, the lower 2 bits of the PCU address bus are not required to decode the address space for the HCU. When reading a register that is less than 32 bits wide zeros should be returned on the upper unused bit(s) of hcu_pcu_datain. The configuration registers of the HCU are listed in Table 193.

TABLE 193 HCU Registers Value Address Register on (HCU_base+) Name #bits Reset Description Control registers 0x00 Reset 1 0x1 A write to this register causes a reset of the HCU. 0x04 Go 1 0x0 Writing 1 to this register starts the HCU. Writing 0 to this register halts the HCU. When Go is asserted all counters, flags etc. are cleared or given their initial value, but configuration registers keep their values. When Go is deasserted the state-machines go to their idle states but all counters and configuration registers keep their values. The HCU should be started after the CFU, SFU, TFU, and DNC. This register can be read to determine if the HCU is running (1 = running, 0 = stopped). Setup registers (constant for during processing) 0x10 AvailMask 4 0x0 Mask used to determine which of the dotgen units etc. are to be checked before a dot is generated by the HCU within the specified margins for the specified color plane. If the specified dotgen unit is stalled, then the HCU will also stall. See Table for bit allocation and definition. 0x14 TMMask 4 0x0 Same as AvailMask, but used in the top margin area before the appropriate target page is reached. 0x18 PageMarginY 32 0x0000_0000 The first line considered to be off the page. 0x1C MaxDot 16 0x0000 This is the maximum dot number − 1 present across a page. For example if a page contains 13824 dots, then MaxDot will be 13823. 0x20 TopMargin 32 0x0000_0000 The first line on a page to be considered within the target page for contone and spot data. (0 = first printed line of page) 0x24 BottomMargin 32 0x0000_0000 The first line in the target bottom margin for contone and spot data (i.e. first line after target page). 0x28 LeftMargin 16 0x0000 The first dot on a line within the target page for contone and spot data. 0x2C RightMargin 16 0xFFFF The first dot on a line within the target right margin for contone and spot data. 0x30 TagTopMargin 32 0x0000_0000 The first line on a page to be considered within the target page for tag data. 0 = first printed line of page) 0x34 TagBottomMargin 32 0x0000_0000 The first line in the target bottom margin for tag data (i.e. first line after target page). 0x38 TagLeftMargin 16 0x0000 The first dot on a line within the target page for tag data. 0x3C TagRightMargin 16 0xFFFF The first dot on a line within the target right margin for tag data. 0x44 StartDMAdr[21:5] 17 0x0_0000 Points to the first 256-bit word of the first line of the dither matrix in DRAM. 0x48 EndDMAdr[21:5] 17 0x0_0000 Points to the last address of the group of four 256- bit reads (or 8 if single buffering) that reads in the last line of the dither matrix. 0x4C LineIncrement 5 0x2 The number of 256-bit words in DRAM from the start of one line of the dither matrix and the start of the next line, i.e. the value by which the DRAM address is incremented at the start of a line so that it points to the start of the next line of the dither matrix. 0x50 DMInitIndexC0 8 0x00 If using the single-buffer scheme this register represents the initial index within 256-byte dither matrix line buffer for contone plane 0. If using double-buffer scheme, only the 7 lsbs are used. 0x54 DMLwrIndexC0 8 0x00 If using the single-buffer scheme this register represents the lower index within 256-byte dither matrix line buffer for contone plane 0. If using double-buffer scheme, only the 7 lsbs are used. 0x58 DMUprIndexC0 8 0x3F If using the single-buffer scheme this register represents the upper index within 256-byte dither matrix line buffer for contone plane 0. After reading the data at this location the index wraps to DMLwrIndexC0. If using double-buffer scheme, only the 7 lsbs are used. 0x5C DMInitIndexC1 8 0x00 If using the single-buffer scheme this register represents the initial index within 256-byte dither matrix line buffer for contone plane 1. If using double-buffer scheme, only the 7 lsbs are used. 0x60 DMLwrIndexC1 8 0x00 If using the single-buffer scheme this register represents the lower index within 256-byte dither matrix line buffer for contone plane 1. If using double-buffer scheme, only the 7 lsbs are used. 0x64 DMUprIndexC1 8 0x3F If using the single-buffer scheme this register represents the upper index within 256-byte dither matrix line buffer for contone plane 1. After reading the data at this location the index wraps to DMLwrIndexC1. If using double-buffer scheme, only the 7 lsbs are used. 0x68 DMInitIndexC2 8 0x00 If using the single-buffer scheme this register represents the initial index within 256-byte dither matrix line buffer for contone plane 2. If using double-buffer scheme, only the 7 lsbs are used. 0x6C DMLwrIndexC2 8 0x00 If using the single-buffer scheme this register represents the lower index within 256-byte dither matrix line buffer for contone plane 2. If using double-buffer scheme, only the 7 lsbs are used. 0x70 DMUprIndexC2 8 0x3F If using the single-buffer scheme this register represents the upper index within 256-byte dither matrix line buffer for contone plane 2. After reading the data at this location the index wraps to DMLwrIndexC2. If using double-buffer scheme, only the 7 lsbs are used. 0x74 DMInitIndexC3 8 0x00 If using the single-buffer scheme this register represents the initial index within 256-byte dither matrix line buffer for contone plane 3. If using double-buffer scheme, only the 7 lsbs are used. 0x78 DMLwrIndexC3 8 0x00 If using the single-buffer scheme this register represents the lower index within 256-byte dither matrix line buffer for contone plane 3. If using double-buffer scheme, only the 7 lsbs are used. 0x7C DMUprIndexC3 8 0x3F If using the single-buffer scheme this register represents the upper index within 256-byte dither matrix line buffer for contone plane 3. After reading the data at this location the index wraps to DMLwrIndexC3. If using double-buffer scheme, only the 7 lsbs are used. 0x80 DoubleLineBuf 1 0x1 Selects the dither line buffer mode to be single or double buffer. 0 - single line buffer mode 1 - double line buffer mode 0x84 to 0x98 IOMappingLo 6 × 32 0x0000_0000 The dot reorg mapping for output inks 0 to 5. For each ink's 64-bit IOMapping value, IOMappingLo represents the low order 32 bits. 0x9C to 0xB0 IOMappingHi 6 × 32 0x0000_0000 The dot reorg mapping for output inks 0 to 5. For each ink's 64-bit IOMapping value, IOMappingHi represents the high order 32 bits. 0xB4 to 0xC0 cpConstant 4 × 8  0x00 The constant contone value to output for contone plane N when printing in the margin areas of the page. This value will typically be 0. 0xC4 sConstant 1 0x0 The constant bi-level value to output for spot when printing in the margin areas of the page. This value will typically be 0. 0xC8 tConstant 1 0x0 The constant bi-level value to output for tag data when printing in the margin areas of the page. This value will typically be 0. 0xCC DitherConstant 8 0xFF The constant value to use for dither matrix when the dither matrix is not available, i.e. when the signal dm_avail is 0. This value will typically be 0xFF so that cpConstant can easily be 0x00 or 0xFF without requiring a dither matrix (DitherConstant is primarily used for threshold dithering in the margin areas). Debug registers (read only) 0xD0 HcuPortsDebug 14 N/A Bit 13 = tfu_hcu_avail Bit 12 = hcu_tfu_advdot Bit 11 = sfu_hcu_avail Bit 10 = hcu_sfu_advdot Bit 9 = cfu_hcu_avail Bit 8 = hcu_cfu_advdot Bit 7 = dnc_hcu_ready Bit 6 = hcu_dnc_avail Bits 5-0 = hcu_dnc_data 0xD4 HcuDotgenDebug 15 N/A Bit 14 = after_top_margin Bit 13 = in_tag_target_page Bit 12 = in_target_page Bit 11 = tp_avail Bit 10 = s_avail Bit 9 = cp_avail Bit 8 = dm_avail Bit 7 = advdot Bits 5-0 = [tp, s, cp3, cp2, cp1, cp0] (i.e. 6 bit input to dot reorg units) 0xD8 HcuDitherDebug1 17 N/A Bit 17 = advdot Bit 16 = dm_avail Bit 15-8 = cp1_dither_val Bits 7-0 = cp0_dither_val 0xDC HcuDitherDebug2 17 N/A Bit 17 = advdot Bit 16 = dm_avail Bit 15-8 = cp3_dither_val Bits 7-0 = cp2_dither_vall 28.4.3 Control Unit

The control unit is responsible for controlling the overall flow of the HCU. It is responsible for determining whether or not a dot will be generated in a given cycle, and what dot will actually be generated—including whether or not the dot is in a margin area, and what dither cell values should be used at the specific dot location. A block diagram of the control unit is shown in FIG. 232.

The inputs to the control unit are a number of avail flags specifying whether or not a given dotgen unit is capable of supplying ‘real’ data in this cycle. The term ‘real’ refers to data generated from external sources, such as contone line buffers, bi-level line buffers, and tag plane buffers. Each dotgen unit informs the control unit whether or not a dot can be generated this cycle from real data. It must also check that the DNC is ready to receive data.

The contone/spot margin unit is responsible for determining whether the current dot coordinate is within the target contone/spot margins, and the tag margin unit is responsible for determining whether the current dot coordinate is within the target tag margins.

The dither matrix table interface provides the interface to DRAM for the generation of dither cell values that are used in the halftoning process in the contone dotgen unit.

28.4.3.1 Determine advdot

The HCU does not always require contone planes, bi-level or tag planes in order to produce a page. For example, a given page may not have a bi-level layer, or a tag layer. In addition, the contone and bi-level parts of a page are only required within the contone and bi-level page margins, and the tag part of a page is only required within the tag page margins. Thus output dots can be generated without contone, bi-level or tag data before the respective top margins of a page has been reached, and 0s are generated for all color planes after the end of the page has been reached (to allow later stages of the printing pipeline to flush).

Consequently the HCU has an AvailMask register that determines which of the various input avail flags should be taken notice of during the production of a page from the first line of the target page, and a TMMask register that has the same behaviour, but is used in the lines before the target page has been reached (i.e. inside the target top margin area). The dither matrix mask bit TMask[0] is the exception, it applies to all margins areas not just the top margin. Each bit in the AvailMask refers to a particular avail bit: if the bit in the AvailMask register is set, then the corresponding avail bit must be 1 for the HCU to advance a dot. The bit to avail correspondence is shown in Table 194. Care should be taken with TMMask—if the particular data is not available after the top margin has been reached, then the HCU will stall. Note that the avail bits for contone and spot colors are ANDed with in_target_page after the target page area has been reached to allow dot production in the contone/spot margin areas without needing any data in the CFU and SFU. The avail bit for tag color is ANDed with in_tag_target_page after the target tag page area has been reached to allow dot production in the tag margin areas without needing any data in the TFU.

TABLE 194 Correspondence between bit in AvailMask and avail flag bit # in AvailMask avail flag description 0 dm_avail dither matrix data available 1 cp_avail contone pixels available 2 s_avail spot color available 3 tp_avail tag plane available

Each of the input avail bits is processed with its appropriate mask bit and the after_top_margin flag (note the dither matrix is the exception it is processed with in_target_page). The output bits are ANDed together along with Go and output_buff_full (which specifies whether the output buffer is ready to receive a dot in this cycle) to form the output bit advdot. We also generate wr_advdot. In this way, if the output buffer is full or any of the specified avail flags is clear, the HCU will stall. When the end of the page is reached, in_page will be deasserted and the HCU will continue to produce 0 for all dots as long as the DNC requests data. A block diagram of the determine advdot unit is shown in FIG. 233.

The advance dot block also determines if current page needs dither matrix, it indicates to the dither matrix table interface block via the dm_read_enable signal. If no dither is required in the margins or in the target page then dm_read_enable will be 0 and no dither will be read in for this page.

28.4.3.2 Position Unit

The position unit is responsible for outputting the position of the current dot (curr_pos, curr_line) and whether or not this dot is the last dot of a line (advline). Both curr_pos and curr_line are set to 0 at reset or when Go transitions from 0 to 1. The position unit relies on the advdot input signal to advance through the dots on a page. Whenever an advdot pulse is received, curr_pos gets incremented. If curr_pos equals max_dot then an advline pulse is generated as this is the last dot in a line, curr_line gets incremented, and the curr_pos is reset to 0 to start counting the dots for the next line.

The position unit also generates a filtered version of advline called dm_advline to indicate to the dither matrix pointers to increment to the next line. The dm_advline is only incremented when dither is required for that line.

if ((after_top_margin AND avail_mask[0]) OR tm_mask[0]) then dm_advline = advline else dm_advline = 0 28.4.3.3 Margin Unit

The responsibility of the margin unit is to determine whether the specific dot coordinate is within the page at all, within the target page or in a margin area (see FIG. 234). This unit is instantiated for both the contone/spot margin unit and the tag margin unit.

The margin unit takes the current dot and line position, and returns three flags.

-   -   the first, in_page is 1 if the current dot is within the page,         and 0 if it is outside the page.     -   the second flag, in_target_page, is 1 if the dot coordinate is         within the target page area of the page, and 0 if it is within         the target top/left/bottom/right margins.     -   the third flag, after_top_margin, is 1 if the current dot is         below the target top margin, and 0 if it is within the target         top margin.

A block diagram of the margin unit is shown in FIG. 235.

28.4.3.4 Dither Matrix Table Interface

The dither matrix table interface provides the interface to DRAM for the generation of dither cell values that are used in the halftoning process in the contone dotgen unit. The control flag dm_read_enable enables the reading of the dither matrix table line structure from DRAM. If dm_read_enable is 0, the dither matrix is not specified in DRAM and no DRAM accesses are attempted. The dither matrix table interface has an output flag dm_avail which specifies if the current line of the specified matrix is available. The HCU can be directed to stall when dm_avail is 0 by setting the appropriate bit in the HCU's AvailMask or TMMask registers. When dm_avail is 0 the value in the DitherConstant register is used as the dither cell values that are output to the contone dotgen unit.

The dither matrix table interface consists of a state machine that interfaces to the DRAM interface, a dither matrix buffer that provides dither matrix values, and a unit to generate the addresses for reading the buffer. FIG. 236 shows a block diagram of the dither matrix table interface.

28.4.3.5 Dither Data Structure in DRAM

The dither matrix is stored in DRAM in 256-bit words, transferred to the HCU in 64-bit words and consumed by the HCU in bytes. Table 195 shows the 64-bit words mapping to 256-bit word addresses, and Table 196 shows the 8-bits dither value mapping in the 64-bits word.

TABLE 195 Dither Data stored in DRAM Address[21:5] Data[255:0] 00000 D3 D2 D1 D0 [255:192] [191:128] [127:64] [63:0] 00001 D7 D6 D5 D4 [255:192] [191:128] [127:64] [63:0] 00010 D11 D10 D9 D8 [255:192] [191:128] [127:64] [63:0] 00011 D15 D14 D13 D12 [255:192] [191:128] [127:64] [63:0] 00100 D19 D18 D17 D16 [255:192] [191:128] [127:64] [63:0] etc

When the HCU first requests data from DRAM, the 64-bits word transfer order will be D0,D1,D2,D3. On the second request the transfer order will be D4,D5,D6,D7 and so on for other requests.

TABLE 196 Dither data stored in HCUs line buffer Dither index [7:0] Data[7:0] 00 D0[7:0] 01 D0[15:8] 02 D0[23:16] 03 D0[31:24] 04 D0[39:32] 05 D0[47:40] 06 D0[55:48] 07 D0[63:56] 08 D1[7:0] 09 D1[15:8] 0A D1[23:16] 0B D1[31:24] 0C D1[39:32] 0D D1[47:40] 0E D1[55:48] 0F D1[63:56] 10 D2[7:0] 11 D2[15:8] 12 D2[23:16] 13 D2[32:24] 14 D2[39:32] 15 D2[47:40] 16 D2[55:48] 17 D2[63:56] 18 D3[7:0] 19 D3[15:8] 1A D3[23:16] 1B D3[31:24] 1C D3[39:32] 1D D3[47:40] 1E D3[55:48] 1F D3[63:56] 20 D4[7:0] 21 D4[15:8] 22 D4[23:16] 23 D4[31:24] 24 D4[39:32] 25 D4[47:40] 26 D4[55:48] 27 D4[63:56] 28 D5[7:0] 29 D5[15:8] 2A D5[23:16] 2B D5[31:24] 2C D5[39:32] 2D D5[47:40] 2E D5[55:48] 2F D5[63:56] etc. etc. 26.4.3.5.1 Dither Matrix Buffer

The state machine loads dither matrix table data a line at a time from DRAM and stores it in a buffer. A single line of the dither matrix is either 256 or 128 8-bit entries, depending on the programmable bit DoubleLineBuf. If this bit is enabled, a double-buffer mechanism is employed such that while one buffer is read from for the current line's dither matrix data (8 bits representing a single dither matrix entry), the other buffer is being written to with the next line's dither matrix data (64-bits at a time). Alternatively, the single buffer scheme can be used, where the data must be loaded at the end of the line, thus incurring a delay.

The single/double buffer is implemented using a 256 byte 3-port register array, two reads, one write port, with the reads clocked at double the system clock rate (320 MHz) allowing 4 reads per clock cycle.

The dither matrix buffer unit also provides the mechanism for keeping track of the current read and write buffers, and providing the mechanism such that a buffer cannot be read from until it has been written to. In this case, each buffer is a line of the dither matrix, i.e. 256 or 128 bytes.

The dither matrix buffer maintains a read and write pointer for the dither matrix. The output value dm_avail is derived by comparing the read and write pointers to determine when the dither matrix is not empty. The write pointer wr_adr is incremented each time a 64-bit word is written to the dither matrix buffer and the read pointer rd_ptr is incremented each time dm_advline is received. If double_line_buf is 0 the rd_ptr will increment by 2, otherwise it will increment by 1. If the dither matrix buffer is full then no further writes will be allowed (buff_full=1), or if the buffer is empty no further buffer reads are allowed (buff_emp=1).

The read addresses are byte aligned and are generated by the read address generator. A single dither matrix entry is represented by 8 bits and an entry is read for each of the four contone planes in parallel. If double buffer is used (double_line_buf=1) the read address is derived from 7-bit address from the read address generator and 1-bit from the read pointer. If double_line_buf=0 then the read address is the full 8-bits from the read address generator.

if (double_line_buf == 1 )then read_port[7:0] = {rd_ptr[0],rd_adr[6:0]} // concatenation else read_port[7:0] = rd_adr[7:0] 28.4.3.5.2 Read Address Generator

For each contone plane there is a initial, lower and upper index to be used when reading dither cell values from the dither matrix double buffer. The read address for each plane is used to select a byte from the current 256-byte read buffer. When Go gets set (0 to 1 transition), or at the end of a line, the read addresses are set to their corresponding initial index. Otherwise, the read address generator relies on advdot to advance the addresses within the inclusive range specified the lower and upper indices, represented by the following pseudocode:

if (advdot == 1) then if (advline == 1) then rd_adr = dm_init_index elsif (rd_adr == dm_upr_index) then rd_adr = dm_lwr_index else rd_adr ++ else rd_adr = rd_adr 28.4.3.5.3 State Machine

The dither matrix is read from DRAM in single 256-bit accesses, receiving the data from the DIU over 4 clock cycles (64-bits per cycle). The protocol and timing for read accesses to DRAM is described in section 20.9.1 on page 240. Read accesses to DRAM are implemented by means of the state machine described in FIG. 238.

All counters and flags should be cleared after reset or when Go transitions from 0 to 1. While the Go bit is 1, the state machine relies on the dm_read_enable bit to tell it whether to attempt to read dither matrix data from DRAM. When dm_read_enable is clear, the state machine does nothing and remains in the idle state. When dm_read_enable is set, the state machine continues to load dither matrix data, 256-bits at a time (received over 4 clock cycles, 64 bits per cycle), while there is space available in the dither matrix buffer, (buff_full !=1).

The read address and line_start_adr are initially set to start_dm_adr. The read address gets incremented after each read access. It takes 4 or 8 read accesses to load a line of dither matrix into the dither matrix buffer, depending on whether we're using a single or double buffer. A count is kept of the accesses to DRAM. When a read access completes and access_count equals 3 or 7, a line of dither matrix has just been loaded from and the read address is updated to line_start_adr plus line_increment so it points to the start of the next line of dither matrix. (line_start_adr is also updated to this value). If the read address equals end_dm_adr then the next read address will be start_dm_adr, thus the read address wraps to point to the start of the area in DRAM where the dither matrix is stored.

The write address for the dither matrix buffer is implemented by means of a modulo-32 counter that is initially set to 0 and incremented when diu_hcu_rvalid is asserted.

FIG. 237 shows an example of setting start_dm_adr and end_dm_adr values in relation to the line increment and double line buffer settings. The calculation of end_dm_adr is

// end_dm_adr calculation dm_height = Dither matrix height in lines if (double_line_buf == 1) // end_dm_adr[21:5] = start_dm_adr[21:5] + (((dm_height − 1)*line_inc) + 3) << 5) else end_dm_adr[21:5] = start_dm_adr[21:5] + (((dm_height − 1)*line_inc) + 7) << 5) 28.4.4 Contone Dotgen Unit

The contone dotgen unit is responsible for producing a dot in up to 4 color planes per cycle. The contone dotgen unit also produces a cp_avail flag which specifies whether or not contone pixels are currently available, and the output hcu_cfu_advdot to request the CFU to provide the next contone pixel in up to 4 color planes.

The block diagram for the contone dotgen unit is shown in FIG. 239.

A dither unit provides the functionality for dithering a single contone plane. The contone image is only defined within the contone/spot margin area. As a result, if the input flag in_target_page is 0, then a constant contone pixel value is used for the pixel instead of the contone plane.

The resultant contone pixel is then halftoned. The dither value to be used in the halftoning process is provided by the control data unit. The halftoning process involves a comparison between a pixel value and its corresponding dither value. If the 8-bit contone value is greater than or equal to the 8-bit dither matrix value a 1 is output. If not, then a 0 is output. This means each entry in the dither matrix is in the range 1-255 (0 is not used).

Note that constant use is dependant on the in_target_page signal only, if in_target_page is 1 then the cfu_hcu_c*_data should be allowed to pass through, regardless of the stalling behaviour or the avail_mask[1] setting. This allows a constant value to be setup on the CFU output data, and the use of different constants while inside and outside the target page. The hcu_cfu_advdot will always be zero if the avail_mask[1] is zero.

28.4.5 Spot Dotgen Unit

The spot dotgen unit is responsible for producing a dot of bi-level data per cycle. It deals with bi-level data (and therefore does not need to halftone) that comes from the LBD via the SFU. Like the contone layer, the bi-level spot layer is only defined within the contone/spot margin area. As a result, if input flag in_target_page is 0, then a constant dot value (typically this would be 0) is used for the output dot.

The spot dotgen unit also produces a s_avail flag which specifies whether or not spot dots are currently available for this spot plane, and the output hcu_sfu_advdot to request the SFU to provide the next bi-level data value. The spot dotgen unit can be represented by the following pseudocode:

s_avail = sfu_hcu_avail if (in_target_page == 1 AND avail_mask[2] == 0 )OR (in_target_page == 0) then hcu_sfu_advdot = 0 else hcu_sfu_advdot = advdot if (in_target_page == 1) then sp = sfu_hcu_sdata else sp = sp_constant

Note that constant use is dependant on the in target_page signal only, if in_target_page is 1 then the sfu_hcu_data should be allowed to pass through, regardless of the stalling behaviour or the avail_mask setting. This allows a constant value to be setup on the SFU output data, and the use of different constants while inside and outside the target page. The hcu_sfu_advdot will always be zero if the avail_mask[2] is zero.

28.4.6 Tag Dotgen Unit

This unit is very similar to the spot dotgen unit (see Section 28.4.5) in that it deals with bi-level data, in this case from the TE via the TFU. The tag layer is only defined within the tag margin area. As a result, if input flag in_tag_target_page is 0, then a constant dot value, tp_constant (typically this would be 0), is used for the output dot. The tagplane dotgen unit also produces a tp_avail flag which specifies whether or not tag dots are currently available for the tagplane, and the output hcu_tfu_advdot to request the TFU to provide the next bi-level data value.

The hcu_tfu_advdot generation is similar to the SFU and CFU, except it depends only on in_target_page and advdot. It does not take into account the avail mask when inside the target page.

28.4.7 Dot Reorg Unit

The dot reorg unit provides a means of mapping the bi-level dithered data, the spot0 color, and the tag data to output inks in the actual printhead. Each dot reorg unit takes a set of 6 1-bit inputs and produces a single bit output that represents the output dot for that color plane.

The output bit is a logical combination of any or all of the input bits. This allows the spot color to be placed in any output color plane (including infrared for testing purposes), black to be merged into cyan, magenta and yellow (in the case of no black ink in the Memjet printhead), and tag dot data to be placed in a visible plane. An output for fixative can readily be generated by simply combining desired input bits.

The dot reorg unit contains a 64-bit lookup to allow complete freedom with regards to mapping. Since all possible combinations of input bits are accounted for in the 64 bit lookup, a given dot reorg unit can take the mapping of other reorg units into account. For example, a black plane reorg unit may produce a 1 only if the contone plane 3 or spot color inputs are set (this effectively composites black bi-level over the contone). A fixative reorg unit may generate a 1 if any 2 of the output color planes is set (taking into account the mappings produced by the other reorg units). If dead nozzle replacement is to be used (see section 29.4.2 on page 473), the dot reorg can be programmed to direct the dots of the specified color into the main plane, and 0 into the other. If a nozzle is then marked as dead in the DNC, swapping the bits between the planes will result in 0 in the dead nozzle, and the required data in the other plane.

If dead nozzle replacement is to be used, and there are no tags, the TE can be programmed with the position of dead nozzles and the resultant pattern used to direct dots into the specified nozzle row. If only fixed background TFS is to be used, a limited number of nozzles can be replaced. If variable tag data is to be used to specify dead nozzles, then large numbers of dead nozzles can be readily compensated for.

The dot reorg unit can be used to average out the nozzle usage when two rows of nozzles share the same ink and tag encoding is not being used. The TE can be programmed to produce a regular pattern (e.g. 0101 on one line, and 1010 on the next) and this pattern can be used as a directive as to direct dots into the specified nozzle row.

Each reorg unit contains a 64-bit IOMapping value programmable as two 32-bit HCU registers, and a set of selection logic based on the 6-bit dot input (2⁶=64 bits), as shown in FIG. 240. The mapping of input bits to each of the 6 selection bits is as defined in Table 197.

TABLE 197 Mapping of input bits to 6 selection bits address bit likely of lookup tied to interpretation 0 bi-level dot from contone layer 0 cyan 1 bi-level dot from contone layer 1 magenta 2 bi-level dot from contone layer 2 yellow 3 bi-level dot from contone layer 3 black 4 bi-level spot0 dot black 5 bi-level tag dot infra-red 28.4.8 Output Buffer

The output buffer de-couples the stalling behaviour of the feeder units from the stalling behaviour of the DNC. The larger the buffer the greater de-coupling. Currently the output buffer size is 2, but could be increased if needed at the cost of extra area.

If the Go bit is set to 0 no read or write of the output buffer is permitted. On a low to high transition of the Go bit the contents of the output buffer are cleared.

The output buffer also implements the interface logic to the DNC. If there is data in the output buffer the hcu_dnc_avail signal will be 1, otherwise is will be 0. If both hcu_dnc_avail and dnc_hcu_ready are 1 then data is read from the output buffer.

On the write side if there is space available in the output buffer the logic indicates to the control unit via the output_buff_full signal. The control unit will then allow writes to the output buffer via the wr_advdot signal. If the writes to the output buffer are after the end of a page (indicated by in_page equal to 0) then all dots written into the output buffer are set to zero.

28.4.8.1 HCU to DNC Interface

FIG. 241 shows the timing diagram and representative logic of the HCU to DNC interface. The hcu_dnc_avail signal indicate to the DNC that the HCU has data available. The dnc_hcu_ready signal indicates to the HCU that the DNC is ready to accept data. When both signals are high data is transferred from the HCU to the DNC. Once the HCU indicates it has data available (setting the hcu_dnc_avail signal high) it can only set the hcu_dnc_avail low again after a dot is accepted by the DNC.

28.4.9 Feeder to HCU Interfaces

FIG. 242 shows the feeder unit to HCU interface timing diagram, and FIG. 243 shows representative logic of the interface with the register positions. sfu_hcu_data and sfu_hcu_avail are always registered while the sfu_hcu_advdot is not. The hcu_sfu_avail signal indicates to the HCU that the feeder unit has data available, and sfu_hcu_advdot indicates to the feeder unit that the HCU has captured the last dot. The HCU can never produce an advance dot pulse while the avail is low. The diagrams show the example of the SFU to HCU interface, but the same interface is used for the other feeder units TFU and CFU.

29 Dead Nozzle Compensator (DNC)

29.1 Overview

The Dead Nozzle Compensator (DNC) is responsible for adjusting Memjet dot data to take account of non-functioning nozzles in the Memjet printhead. Input dot data is supplied from the HCU, and the corrected dot data is passed out to the DWU. The high level data path is shown by the block diagram in FIG. 244.

The DNC compensates for a dead nozzles by performing the following operations:

-   -   Dead nozzle removal, i.e. turn the nozzle off     -   Ink replacement by direct substitution i.e. K->K     -   Ink replacement by indirect substitution i.e. K->CMY     -   Error diffusion to adjacent nozzles     -   Fixative corrections

The DNC is required to efficiently support up to 5% dead nozzles, under the expected DRAM bandwidth allocation, with no restriction on where dead nozzles are located and handle any fixative correction due to nozzle compensations. Performance must degrade gracefully after 5% dead nozzles.

29.2 Dead Nozzle Identification

Dead nozzles are identified by means of a position value and a mask value. Position information is represented by a 10-bit delta encoded format, where the 10-bit value defines the number of dots between dead nozzle columns¹⁹. With the delta information it also reads the 6-bit dead nozzle mask (dn_mask) for the defined dead nozzle position. Each bit in the dn_mask corresponds to an ink plane. A set bit indicates that the nozzle for the corresponding ink plane is dead. The dead nozzle table format is shown in FIG. 245. The DNC reads dead nozzle information from DRAM in single 256-bit accesses. A 10-bit delta encoding scheme is chosen so that each table entry is 16 bits wide, and 16 entries fit exactly in each 256-bit read. Using 10-bit delta encoding means that the maximum distance between dead nozzle columns is 1023 dots. It is possible that dead nozzles may be spaced further than 1023 dots from each other, so a null dead nozzle identifier is required. A null dead nozzle identifier is defined as a 6-bit dn_mask of all zeros. These null dead nozzle identifiers should also be used so that: ¹⁹for a 10-bit delta value of d, if the current column n is a dead nozzle column then the next dead nozzle column is given by n+(d+1).

-   -   the dead nozzle table is a multiple of 16 entries (so that it is         aligned to the 256-bit DRAM locations)     -   the dead nozzle table spans the complete length of the line,         i.e. the first entry dead nozzle table should have a delta from         the first nozzle column in a line and the last entry in the dead         nozzle table should correspond to the last nozzle column in a         line.

Note that the DNC deals with the width of a page. This may or may not be the same as the width of the printhead (the PHI may introduce some margining to the page so that its dot output matches the width of the printhead). Care must be taken when programming the dead nozzle table so that dead nozzle positions are correctly specified with respect to the page and printhead.

29.3 Dram Storage and Bandwidth Requirement

The memory required is largely a factor of the number of dead nozzles present in the printhead (which in turn is a factor of the printhead size). The DNC is required to read a 16-bit entry from the dead nozzle table for every dead nozzle. Table 198 shows the DRAM storage and average²⁰ bandwidth requirements for the DNC for different percentages of dead nozzles and different page sizes. ²⁰Average bandwidth assumes an even spread of dead nozzles. Clumps of dead nozzles may cause delays due to insufficient available DRAM bandwidth. These delays will occur every line causing an accumulative delay over a page.

TABLE 198 Dead Nozzle storage and average bandwidth requirements Dead nozzle table % Dead Memory Bandwidth Page size Nozzles (KBytes) (bits/cycle) A4^(a)  5% 1.4^(c) 0.8^(d) 10% 2.7 1.6 15% 4.1 2.4 A3^(b)  5% 1.9 0.8 10% 3.8 1.6 15% 5.7 2.4 ^(a)Bi-lithic printhead has 13824 nozzles per color providing full bleed printing for A4/Letter ^(b)Bi-lithic printhead has 19488 nozzles per color providing full bleed printing for A3 ^(c)16 bits × 13824 nozzles × 0.05 dead ^(d)(16 bits read/20 cycles) = 0.8 bits/cycle 29.4 Nozzle Compensation

DNC receives 6 bits of dot information every cycle from the HCU, 1 bit per color plane. When the dot position corresponds to a dead nozzle column, the associated 6-bit dn_mask indicates which ink plane(s) contains a dead nozzle(s). The DNC first deletes dots destined for the dead nozzle. It then replaces those dead dots, either by placing the data destined for the dead nozzle into an adjacent ink plane (direct substitution) or into a number of ink planes (indirect substitution). After ink replacement, if a dead nozzle is made active again then the DNC performs error diffusion. Finally, following the dead nozzle compensation mechanisms the fixative, if present, may need to be adjusted due to new nozzles being activated, or dead nozzles being removed.

29.4.1 Dead Nozzle Removal

If a nozzle is defined as dead, then the first action for the DNC is to turn off (zeroing) the dot data destined for that nozzle. This is done by a bit-wise ANDing of the inverse of the dn_mask with the dot value.

29.4.2 Ink Replacement

Ink replacement is a mechanism where data destined for the dead nozzle is placed into an adjacent ink plane of the same color (direct substitution, i.e. K->K_(alternative)), or placed into a number of ink planes, the combination of which produces the desired color (indirect substitution, i.e. K->CMY). Ink replacement is performed by filtering out ink belonging to nozzles that are dead and then adding back in an appropriately calculated pattern. This two step process allows the optional re-inclusion of the ink data into the original dead nozzle position to be subsequently error diffused. In the general case, fixative data destined for a dead nozzle should not be left active intending it to be later diffused.

The ink replacement mechanism has 6 ink replacement patterns, one per ink plane, programmable by the CPU. The dead nozzle mask is ANDed with the dot data to see if there are any planes where the dot is active but the corresponding nozzle is dead. The resultant value forms an enable, on a per ink basis, for the ink replacement process. If replacement is enabled for a particular ink, the values from the corresponding replacement pattern register are ORed into the dot data. The output of the ink replacement process is then filtered so that error diffusion is only allowed for the planes in which error diffusion is enabled. The output of the ink replacement logic is ORed with the resultant dot after dead nozzle removal. See Figure n page 565 on page Error! Bookmark not defined. for implementation details.

For example if we consider the printhead color configuration C,M,Y,K₁,K₂,IR and the input dot data from the HCU is b101100. Assuming that the K₁ ink plane and IR ink plane for this position are dead so the dead nozzle mask is b000101. The DNC first removes the dead nozzle by zeroing the K₁ plane to produce b101000. Then the dead nozzle mask is ANDed with the dot data to give b000100 which selects the ink replacement pattern for K₁ (in this case the ink replacement pattern for K₁ is configured as b000010, i.e. ink replacement into the K₂ plane). Providing error diffusion for K₂ is enabled, the output from the ink replacement process is b000010. This is ORed with the output of dead nozzle removal to produce the resultant dot b101010. As can be seen the dot data in the defective K₁ nozzle was removed and replaced by a dot in the adjacent K₂ nozzle in the same dot position, i.e. direct substitution.

In the example above the K₁ ink plane could be compensated for by indirect substitution, in which case ink replacement pattern for K₁ would be configured as b111000 (substitution into the CMY color planes), and this is ORed with the output of dead nozzle removal to produce the resultant dot b111000. Here the dot data in the defective K, ink plane was removed and placed into the CMY ink planes.

29.4.3 Error Diffusion

Based on the programming of the lookup table the dead nozzle may be left active after ink replacement. In such cases the DNC can compensate using error diffusion. Error diffusion is a mechanism where dead nozzle dot data is diffused to adjacent dots.

When a dot is active and its destined nozzle is dead, the DNC will attempt to place the data into an adjacent dot position, if one is inactive. If both dots are inactive then the choice is arbitrary, and is determined by a pseudo random bit generator. If both neighbor dots are already active then the bit cannot be compensated by diffusion.

Since the DNC needs to look at neighboring dots to determine where to place the new bit (if required), the DNC works on a set of 3 dots at a time. For any given set of 3 dots, the first dot received from the HCU is referred to as dot A, and the second as dot B, and the third as dot C. The relationship is shown in FIG. 246.

For any given set of dots ABC, only B can be compensated for by error diffusion if B is defined as dead. A 1 in dot B will be diffused into either dot A or dot C if possible. If there is already a 1 in dot A or dot C then a 1 in dot B cannot be diffused into that dot.

The DNC must support adjacent dead nozzles. Thus if dot A is defined as dead and has previously been compensated for by error diffusion, then the dot data from dot B should not be diffused into dot A. Similarly, if dot C is defined as dead, then dot data from dot B should not be diffused into dot C.

Error diffusion should not cross line boundaries. If dot B contains a dead nozzle and is the first dot in a line then dot A represents the last dot from the previous line. In this case an active bit on a dead nozzle of dot B should not be diffused into dot A. Similarly, if dot B contains a dead nozzle and is the last dot in a line then dot C represents the first dot of the next line. In this case an active bit on a dead nozzle of dot B should not be diffused into dot C.

Thus, as a rule, a 1 in dot B cannot be diffused into dot A if

-   -   a 1 is already present in dot A,     -   dot A is defined as dead,     -   or dot A is the last dot in a line.

Similarly, a 1 in dot B cannot be diffused into dot C if

-   -   a 1 is already present in dot C,     -   dot C is defined as dead,     -   or dot C is the first dot in a line.

If B is defined to be dead and the dot value for B is 0, then no compensation needs to be done and dots A and C do not need to be changed.

If B is defined to be dead and the dot value for B is 1, then B is changed to 0 and the DNC attempts to place the 1 from B into either A or C:

-   -   If the dot can be placed into both A and C, then the DNC must         choose between them. The preference is given by the current         output from the random bit generator, 0 for “prefer left”         (dot A) or 1 for “prefer right” (dot C).     -   If dot can be placed into only one of A and C, then the 1 from B         is placed into that position.     -   If dot cannot be placed into either one of A or C, then the DNC         cannot place the dot in either position.

TABLE 199 Error Diffusion Truth Table when dot B is dead Input A OR A dead C OR C dead OR A last OR C first Output in line B in line Rand‘a A B C 0 0 0 X A input 0 C input 0 0 1 X A input 0 C input 0 1 0 0 1‘b 0 C input 0 1 0 1 A input 0 1 0 1 1 X 1 0 C input 1 0 0 X A input 0 C input 1 0 1 X A input 0 C input 1 1 0 X A input 0 1 1 1 1 X A input 0 C input

Table 199 shows the truth table for DNC error diffusion operation when dot B is defined as dead.

-   a. Output from random bit generator. Determines direction of error     diffusion (0=left, 1=right) -   b. Bold emphasis is used to show the DNC inserted a 1

The random bit value used to arbitrarily select the direction of diffusion is generated by a 32-bit maximum length random bit generator. The generator generates a new bit for each dot in a line regardless of whether the dot is dead or not. The random bit generator can be initialized with a 32-bit programmable seed value.

29.4.4 Fixative Correction

After the dead nozzle compensation methods have been applied to the dot data, the fixative, if present, may need to be adjusted due to new nozzles being activated, or dead nozzles being removed. For each output dot the DNC determines if fixative is required (using the FixativeRequiredMask register) for the new compensated dot data word and whether fixative is activated already for that dot. For the DNC to do so it needs to know the color plane that has fixative, this is specified by the FixativeMask1 configuration register. Table 200 indicates the actions to take based on these calculations.

TABLE 200 Truth table for fixative correction Fixative Present Fixative required Action 1 1 Output dot as is. 1 0 Clear fixative plane. 0 1 Attempt to add fixative. 0 0 Output dot as is.

The DNC also allows the specification of another fixative plane, specified by the FixativeMask2 configuration register, with FixativeMask1 having the higher priority over FixativeMask2. When attempting to add fixative the DNC first tries to add it into the planes defined by FixativeMask1. However, if any of these planes is dead then it tries to add fixative by placing it into the planes defined by FixativeMask2.

Note that the fixative defined by FixativeMask1 and FixativeMask2 could possibly be multi-part fixative, i.e. 2 bits could be set in FixativeMask1 with the fixative being a combination of both inks.

29.5 Implementation

A block diagram of the DNC is shown in FIG. 247.

29.5.1 Definitions of I/O

TABLE 201 DNC port list and description Port name Pins I/O Description Clocks and Resets Pclk 1 In System Clock. prst_n 1 In System reset, synchronous active low. PCU interface pcu_dnc_sel 1 In Block select from the PCU. When pcu_dnc_sel is high both pcu_adr and pcu_dataout are valid. pcu_rwn 1 In Common read/not-write signal from the PCU. pcu_adr[6:2] 5 In PCU address bus. Only 5 bits are required to decode the address space for this block. pcu_dataout[31:0] 32 In Shared write data bus from the PCU. dnc_pcu_rdy 1 Out Ready signal to the PCU. When dnc_pcu_rdy is high it indicates the last cycle of the access. For a write cycle this means pcu_dataout has been registered by the block and for a read cycle this means the data on dnc_pcu_datain is valid. dnc_pcu_datain[31:0] 32 Out Read data bus to the PCU. DIU interface dnc_diu_rreq 1 Out DNC unit requests DRAM read. A read request must be accompanied by a valid read address. dnc_diu_radr[21:5] 17 Out Read address to DIU, 256-bit word aligned. diu_dnc_rack 1 In Acknowledge from DIU that read request has been accepted and new read address can be placed on dnc_diu_radr diu_dnc_rvalid 1 In Read data valid, active high. Indicates that valid read data is now on the read data bus, diu_data. diu_data[63:0] 64 In Read data from DIU. HCU interface dnc_hcu_ready 1 Out Indicates that DNC is ready to accept data from the HCU. hcu_dnc_avail 1 In Indicates valid data present on hcu_dnc_data. hcu_dnc_data[5:0] 6 In Output bi-level dot data in 6 ink planes. DWU interface dwu_dnc_ready 1 In Indicates that DWU is ready to accept data from the DNC. dnc_dwu_avail 1 Out Indicates valid data present on dnc_dwu_data. dnc_dwu_data[5:0] 6 Out Output bi-level dot data in 6 ink planes. 29.5.2 Configuration Registers

The configuration registers in the DNC are programmed via the PCU interface. Refer to section 21.8.2 on page 321 for the description of the protocol and timing diagrams for reading and writing registers in the DNC. Note that since addresses in SoPEC are byte aligned and the PCU only supports 32-bit register reads and writes, the lower 2 bits of the PCU address bus are not required to decode the address space for the DNC. When reading a register that is less than 32 bits wide zeros should be returned on the upper unused bit(s) of dnc_pcu_datain. Table 202 lists the configuration registers in the DNC.

TABLE 202 DNC configuration registers Value Address Register on (DNC_base+) name #bits reset Description Control registers 0x00 Reset 1 0x1 A write to this register causes a reset of the DNC. 0x04 Go 1 0x0 Writing 1 to this register starts the DNC. Writing 0 to this register halts the DNC. When Go is asserted all counters, flags etc. are cleared or given their initial value, but configuration registers keep their values. When Go is deasserted the state- machines go to their idle states but all counters and configuration registers keep their values. This register can be read to determine if the DNC is running (1 = running, 0 = stopped). Setup registers (constant during processing) 0x10 MaxDot 16 0x0000 This is the maximum dot number − 1 present across a page. For example if a page contains 13824 dots, then MaxDot will be 13823. Note that this number may or may not be the same as the number of dots across the printhead as some margining may be introduced in the PHI. 0x14 LSFR 32 0x0000_0000 The current value of the LFSR register used as the 32-bit maximum length random bit generator. Users can write to this register to program a seed value for the 32-bit maximum length random bit generator. Must not be all 1s for taps implemented in XNOR form. (It is expected that writing a seed value will not occur during the operation of the LFSR). This LSFR value could also have a possible use as a random source in program code. 0x20 FixativeMask1 6 0x00 Defines the higher priority fixative plane(s). Bit 0 represents the settings for plane 0, bit 1 for plane 1 etc. For each bit: 1 = the ink plane contains fixative. 0 = the ink plane does not contain fixative. 0x24 FixativeMask2 6 0x00 Defines the lower priority fixative plane(s). Bit 0 represents the settings for plane 0, bit 1 for plane 1 etc. Used only when FixativeMask1 planes are dead. For each bit: 1 = the ink plane contains fixative. 0 = the ink plane does not contain fixative. 0x28 FixativeRequired 6 0x00 Identifies the ink planes that require Mask fixative. Bit 0 represents the settings for plane 0, bit 1 for plane 1 etc. For each bit: 1 = the ink plane requires fixative. 0 = the ink plane does not require fixative (e.g. ink is self-fixing) 0x30 DnTableStartAdr[21:5] 17 0x0_0000 Start address of Dead Nozzle Table in DRAM, specified in 256-bit words. 0x34 DnTableEndAdr[21:5] 17 0x0_0000 End address of Dead Nozzle Table in DRAM, specified in 256-bit words, i.e. the location containing the last entry in the Dead Nozzle Table. The Dead Nozzle Table should be aligned to a 256-bit boundary, if necessary it can be padded with null entries. 0x40-0x54 PlaneReplacePattern[5:0] 6 × 6 0x00 Defines the ink replacement pattern for each of the 6 ink planes. PlaneReplacePattern[0] is the ink replacement pattern for plane 0, PlaneReplacePattern[1] is the ink replacement pattern for plane 1, etc. For each 6-bit replacement pattern for a plane, a 1 in any bit positions indicates the alternative ink planes to be used for this plane. 0x58 DiffuseEnable 6 0x3F Defines whether, after ink replacement, error diffusion is allowed to be performed on each plane. Bit 0 represents the settings for plane 0, bit 1 for plane 1 etc. For each bit: 1 = error diffusion is enabled 0 = error diffusion is disabled Debug registers (read only) 0x60 DncOutputDebug 8 N/A Bit 7 = dwu_dnc_ready Bit 6 = dnc_dwu_avail Bits 5-0 = dnc_dwu_data 0x64 DncReplaceDebug 14 N/A Bit 13 = edu_ready Bit 12 = iru_avail Bits 11-6 = iru_dn_mask Bits 5-0 = iru_data 0x68 DncDiffuseDebug 14 N/A Bit 13 = dwu_dnc_ready Bit 12 = dnc_dwu_avail Bits 11-6 = edu_dn_mask Bits 5-0 = edu_data 29.5.3 Ink Replacement Unit

FIG. 248 shows a sub-block diagram for the ink replacement unit.

29.5.3.1 Control Unit

The control unit is responsible for reading the dead nozzle table from DRAM and making it available to the DNC via the dead nozzle FIFO. The dead nozzle table is read from DRAM in single 256-bit accesses, receiving the data from the DIU over 4 clock cycles (64-bits per cycle). The protocol and timing for read accesses to DRAM is described in section 20.9.1 on page 240. Reading from DRAM is implemented by means of the state machine shown in FIG. 249. All counters and flags should be cleared after reset. When Go transitions from 0 to 1 all counters and flags should take their initial value. While the Go bit is 1, the state machine requests a read access from the dead nozzle table in DRAM provided there is enough space in its FIFO.

A modulo-4 counter, rd_count, is used to count each of the 64-bits received in a 256-bit read access. It is incremented whenever diu_dnc_rvalid is asserted. When Go is 1, dn_table_radr is set to dn_table_start_adr. As each 64-bit value is returned, indicated by diu_dnc_rvalid being asserted, dn_table_radr is compared to dn_table_end_adr.

-   -   If rd_count equals 3 and dn_table_radr equals dn_table_end_adr,         then dn_table_radr is updated to dn_table_start_adr.     -   If rd_count equals 3 and dn_table_radr does not equal         dn_table_end_adr, then dn_table_radr is incremented by 1.

A count is kept of the number of 64-bit values in the FIFO. When diu_dnc_rvalid is 1 data is written to the FIFO by asserting wr_en, and fifo_contents and fifo_wr_adr are both incremented.

When fifo_contents[3:0] is greater than 0 and edu_ready is 1, dnc_hcu_ready is asserted to indicate that the DNC is ready to accept dots from the HCU. If hcu_dnc_avail is also 1 then a dotadv pulse is sent to the GenMask unit, indicating the DNC has accepted a dot from the HCU, and iru_avail is also asserted. After Go is set, a single preload pulse is sent to the GenMask unit once the FIFO contains data.

When a rd_adv pulse is received from the GenMask unit, fifo_rd_adr[4:0] is then incremented to select the next 16-bit value. If fifo_rd_adr[1:0]=11 then the next 64-bit value is read from the FIFO by asserting rd_en, and fifo_contents[3:0] is decremented.

29.5.3.2 Dead Nozzle FIFO

The dead nozzle FIFO conceptually is a 64-bit input, and 16-bit output FIFO to account for the 64-bit data transfers from the DIU, and the individual 16-bit entries in the dead nozzle table that are used in the GenMask unit. In reality, the FIFO is actually 8 entries deep and 64-bits wide (to accommodate two 256-bit accesses).

On the DRAM side of the FIFO the write address is 64-bit aligned while on the GenMask side the read address is 16-bit aligned, i.e. the upper 3 bits are input as the read address for the FIFO and the lower 2 bits are used to select 16 bits from the 64 bits (1st 16 bits read corresponds to bits 15-0, second 16 bits to bits 31-16 etc.).

29.5.3.3 GenMask Unit

The GenMask unit generates the 6-bit dn_mask that is sent to the replace unit. It consists of a 10-bit delta counter and a mask register.

After Go is set, the GenMask unit will receive a preload pulse from the control unit indicating the first dead nozzle table entry is available at the output of the dead nozzle FIFO and should be loaded into the delta counter and mask register. A rd_adv pulse is generated so that the next dead nozzle table entry is presented at the output of the dead nozzle FIFO. The delta counter is decremented every time a dotadv pulse is received. When the delta counter reaches 0, it gets loaded with the current delta value output from the dead nozzle FIFO, i.e. bits 15-6, and the mask register gets loaded with mask output from the dead nozzle FIFO, i.e. bits 5-0. A rd_adv pulse is then generated so that the next dead nozzle table entry is presented at the output of the dead nozzle FIFO.

When the delta counter is 0 the value in the mask register is output as the dn_mask, otherwise the dn_mask is all 0s.

The GenMask unit has no knowledge of the number of dots in a line, it simply loads a counter to count the delta from one dead nozzle column to the next. Thus as described in section 29.2 on page 472 the dead nozzle table should include null identifiers if necessary so that the dead nozzle table covers the first and last nozzle column in a line.

29.5.3.4 Replace Unit

Dead nozzle removal and ink replacement are implemented by the combinatorial logic shown in FIG. 250. Dead nozzle removal is performed by bit-wise ANDing of the inverse of the dn_mask with the dot value.

The ink replacement mechanism has 6 ink replacement patterns, one per ink plane, programmable by the CPU. The dead nozzle mask is ANDed with the dot data to see if there are any planes where the dot is active but the corresponding nozzle is dead. The resultant value forms an enable, on a per ink basis, for the ink replacement process. If replacement is enabled for a particular ink, the values from the corresponding replacement pattern register are ORed into the dot data. The output of the ink replacement process is then filtered so that error diffusion is only allowed for the planes in which error diffusion is enabled.

The output of the ink replacement process is ORed with the resultant dot after dead nozzle removal. If the dot position does not contain a dead nozzle then the dn_mask will be all 0s and the dot, hcu_dnc_data, will be passed through unchanged.

29.5.4 Error Diffusion Unit

FIG. 251 shows a sub-block diagram for the error diffusion unit.

29.5.4.1 Random Bit Generator

The random bit value used to arbitrarily select the direction of diffusion is generated by a maximum length 32-bit LFSR. The tap points and feedback generation are shown in FIG. 252. The LFSR generates a new bit for each dot in a line regardless of whether the dot is dead or not, i.e shifting of the LFSR is enabled when advdot equals 1. The LFSR can be initialised with a 32-bit programmable seed value, random_seed. This seed value is loaded into the LFSR whenever a write occurs to the RandomSeed register. Note that the seed value must not be all 1 s as this causes the LFSR to lock-up.

29.5.4.2 Advance Dot Unit

The advance dot unit is responsible for determining in a given cycle whether or not the error diffuse unit will accept a dot from the ink replacement unit or make a dot available to the fixative correct unit and on to the DWU. It therefore receives the dwu_dnc_ready control signal from the DWU, the iru_avail flag from the ink replacement unit, and generates dnc_dwu_avail and edu_ready control flags.

Only the dwu_dnc_ready signal needs to be checked to see if a dot can be accepted and asserts edu_ready to indicate this. If the error diffuse unit is ready to accept a dot and the ink replacement unit has a dot available, then a advdot pulse is given to shift the dot into the pipeline in the diffuse unit. Note that since the error diffusion operates on 3 dots, the advance dot unit ignores dwu_dnc_ready initially until 3 dots have been accepted by the diffuse unit. Similarly dnc_dwu_avail is not asserted until the diffuse unit contains 3 dots and the ink replacement unit has a dot available.

29.5.4.3 Diffuse Unit

The diffuse unit contains the combinatorial logic to implement the truth table from Table. The diffuse unit receives a dot consisting of 6 color planes (1 bit per plane) as well as an associated 6-bit dead nozzle mask value.

Error diffusion is applied to all 6 planes of the dot in parallel. Since error diffusion operates on 3 dots, the diffuse unit has a pipeline of 3 dots and their corresponding dead nozzle mask values.

The first dot received is referred to as dot A, and the second as dot B, and the third as dot C. Dots are shifted along the pipeline whenever advdot is 1. A count is also kept of the number of dots received. It is incremented whenever advdot is 1, and wraps to 0 when it reaches max_dot. When the dot count is 0 dot C corresponds to the first dot in a line. When the dot count is 1 dot A corresponds to the last dot in a line.

In any given set of 3 dots only dot B can be defined as containing a dead nozzle(s). Dead nozzles are identified by bits set in iru_dn_mask. If dot B contains a dead nozzle(s), the corresponding bit(s) in dot A, dot C, the dead nozzle mask value for A, the dead nozzle mask value for C, the dot count, as well as the random bit value are input to the truth table logic and the dots A, B and C assigned accordingly. If dot B does not contain a dead nozzle then the dots are shifted along the pipeline unchanged.

29.5.5 Fixative Correction Unit

The fixative correction unit consists of combinatorial logic to implement fixative correction as defined in Table 203. For each output dot the DNC determines if fixative is required for the new compensated dot data word and whether fixative is activated already for that dot.

FixativePresent = ((FixativeMask1 | FixativeMask2) & edu_data) != 0 FixativeRequired = (FixativeRequiredMask & edu_data) != 0

It then looks up the truth table to see what action, if any, needs to be taken.

TABLE 203 Truth table for fixative correction Fixative Fixative Present required Action Output 1 1 Output dot as is. dnc_dwu_data = edu_data 1 0 Clear fixative plane. dnc_dwu_data = (edu_data) & ~(FixativeMask1 | FixativeMask2) 0 1 Attempt to add if (FixativeMask1 & fixative. DnMask) != 0 dnc_dwu_data = (edu_data) | (FixativeMask2 & ~DnMask) else dnc_dwu_data = (edu_data) | (FixativeMask1) 0 0 Output dot as is. dnc_dwu_data = edu_data

When attempting to add fixative the DNC first tries to add it into the plane defined by FixativeMask1. However, if this plane is dead then it tries to add fixative by placing it into the plane defined by FixativeMask2. Note that if both FixativeMask1 and FixativeMask2 are both all 0s then the dot data will not be changed.

30 Dotline Writer Unit (DWU)

30.1 Overview

The Dotline Writer Unit (DWU) receives 1 dot (6 bits) of color information per cycle from the DNC. Dot data received is bundled into 256-bit words and transferred to the DRAM. The DWU (in conjunction with the LLU) implements a dot line FIFO mechanism to compensate for the physical placement of nozzles in a printhead, and provides data rate smoothing to allow for local complexities in the dot data generate pipeline.

30.2 Physical Requirement Imposed by the Printhead

The physical placement of nozzles in the printhead means that in one firing sequence of all nozzles, dots will be produced over several print lines. The printhead consists of 12 rows of nozzles, one for each color of odd and even dots. Odd and even nozzles are separated by D₂ print lines and nozzles of different colors are separated by D₁ print lines. See FIG. 254 for reference. The first color to be printed is the first row of nozzles encountered by the incoming paper. In the example this is color 0 odd, although is dependent on the printhead type (see [10] for other printhead arrangments). Paper passes under printhead moving downwards.

For example if the physical separation of each half row is 80 μm equating to D₁=D₂=5 print lines at 1600 dpi. This means that in one firing sequence, color 0 odd nozzles will fire on dotline L, color 0 even nozzles will fire on dotline L-D₁, color 1 odd nozzles will fire on dotline L-D₁-D₂ and so on over 6 color planes odd and even nozzles. The total number of lines fired over is given as 0+5+5 . . . +5=0+11×5=55. See FIG. 255 for example diagram.

It is expected that the physical spacing of the printhead nozzles will be 80 μm (or 5 dot lines), although there is no dependency on nozzle spacing. The DWU is configurable to allow other line nozzle spacings.

TABLE 204 Relationship between Nozzle color/sense and line firing Even line Odd line encountered first encountered first Color Sense line sense line Color 0 Even L even L-5 Odd L-5 odd L Color 1 Even L-10 even L-15 Odd L-15 odd L-10 Color 2 Even L-20 even L-25 Odd L-25 odd L-20 Color 3 Even L-30 even L-35 Odd L-35 odd L-30 Color 4 Even L-40 even L-45 Odd L-45 odd L-40 Color 5 Even L-50 even L-55 Odd L-55 odd L-50 30.3 Line Rate De-Coupling

The DWU block is required to compensate for the physical spacing between lines of nozzles. It does this by storing dot lines in a FIFO (in DRAM) until such time as they are required by the LLU for dot data transfer to the printhead interface. Colors are stored separately because they are needed at different times by the LLU. The dot line store must store enough lines to compensate for the physical line separation of the printhead but can optionally store more lines to allow system level data rate variation between the read (printhead feed) and write sides (dot data generation pipeline) of the FIFOs.

A logical representation of the FIFOs is shown in FIG. 256, where N is defined as the optional number of extra half lines in the dot line store for data rate de-coupling.

30.4 Dot Line Store Storage Requirements

For an arbitrary page width of d dots (where d is even), the number of dots per half line is d/2.

For interline spacing of D₂ and inter-color spacing of D₁, with C colors of odd and even half lines, the number of half line storage is (C−1)(D₂+D₁)+D1.

For N extra half line stores for each color odd and even, the storage is given by (N*C*2).

The total storage requirement is ((C−1)(D₂+D₁)+D1+(N*C*2))*d/2 in bits.

Note that when determining the storage requirements for the dot line store, the number of dots per line is the page width and not necessarily the printhead width. The page width is often the dot margin number of dots less than the printhead width. They can be the same size for full bleed printing.

For example in an A4 page a line consists of 13824 dots at 1600 dpi, or 6912 dots per half dot line. To store just enough dot lines to account for an inter-line nozzle spacing of 5 dot lines it would take 55 half dot lines for color 5 odd, 50 dot lines for color 5 even and so on, giving 55+50+45 . . . 10+5+0=330 half dot lines in total. If it is assumed that N=4 then the storage required to store 4 extra half lines per color is 4×12=48, in total giving 330+48=378 half dot lines. Each half dot line is 6912 dots, at 1 bit per dot give a total storage requirement of 6912 dots×378 half dot lines/8 bits=Approx 319 Kbytes. Similarly for an A3 size page with 19488 dots per line, 9744 dots per half line×378 half dot lines/8=Approx 899 Kbytes.

TABLE 205 Storage requirement for dot line store Lines Storage Lines Storage Page Nozzle required (N = 0) required (N = 4) size Spacing (N = 0) Kbytes (N = 4) Kbytes A4 4 264 223 312 263 5 330 278 378 319 A3 4 264 628 312 742 5 330 785 378 899

The potential size of the dot line store makes it unfeasible to be implemented in on-chip SRAM, requiring the dot line store to be implemented in embedded DRAM. This allows a configurable dotline store where unused storage can be redistributed for use by other parts of the system.

30.5 Nozzle Row Skew

Due to construction limitations of the bi-lithic printhead it is possible that nozzle rows may be misaligned relative to each other. Odd and even rows, and adjacent color rows may be horizontally misaligned by up to 2 dot positions. Vertical misalignment can also occur but is compensated for in the LLU and not considered here. The DWU is required to compensate for the horizontal misalignment.

Dot data from the HCU (through the DNC) produces a dot of 6 colors all destined for the same physical location on paper. If the nozzle rows in the printhead are aligned as shown in FIG. 254 then no adjustment of the dot data is needed.

A conceptual misaligned printhead is shown in FIG. 257. The exact shape of the row alignment is arbitrary, although is most likely to be sloping (if sloping, it could be sloping in either direction). The DWU is required to adjust the shape of the dot streams to take account of the join between printhead ICs. The introduction of the join shape before the data is written to the DRAM means that the PHI sees a single crossover point in the data since all lines are the same length and the crossover point (since all rows are of equal length) is a vertical line—i.e. the crossover is at the same time for all even rows, and at the same time for all odd rows as shown in FIG. 258.

To insert the shape of the join into the dot stream, for each line we must first insert the dots for non-printable area 1, then the printable area data (from the DNC), and then finally the dots for non-printable area 2. This can also be considered as: first produce the dots for non-printable area 1 for line n, and then a repetition of:

-   -   produce the dots for the printable area for line n (from the         DNC)     -   produce the dots for the non-printable area 2 (for line n)         followed by the dots of non-printable area 1 (for line n+1)

The reason for considering the problem this way is that regardless of the shape of the join, the shape of non-printable area 2 merged with the shape of non-printable area 1 will always be a rectangle since the widths of non-printable areas 1 and 2 are identical and the lengths of each row are identical. Hence step 2 can be accomplished by simply inserting a constant number (MaxNozzleSkew) of 0 dots into the stream.

For example, if the color n even row non-printable area 1 is of length X, then the length of color n even row non-printable area 2 will be of length MaxNozzleSkew−X. The split between non-printable areas 1 and 2 is defined by the NozzleSkew registers.

Data from the DNC is destined for the printable area only, the DWU must generate the data destined for the non-printable areas, and insert DNC dot data correctly into the dot data stream before writing dot data to the fifos. The DWU inserts the shape of the misalignment into the dot stream by delaying dot data destined to different nozzle rows by the relative misalignment skew amount.

30.6 Local Buffering

An embedded DRAM is expected to be of the order of 256 bits wide, which results in 27 words per half line of an A4 page, and 54 words per half line of A3. This requires 27 words×12 half colors (6 colors odd and even)=324×256-bit DRAM accesses over a dotline print time, equating to 6 bits per cycle (equal to DNC generate rate of 6 bits per cycle). Each half color is required to be double buffered, while filling one buffer the other buffer is being written to DRAM. This results in 256 bits×2 buffers×12 half colors i.e. 6144 bits in total.

The buffer requirement can be reduced, by using 1.5 buffering, where the DWU is filling 128 bits while the remaining 256 bits are being written to DRAM. While this reduces the required buffering locally it increases the peak bandwidth requirement to the DRAM. With 2× buffering the average and peak DRAM bandwidth requirement is the same and is 6 bits per cycle, alternatively with 1.5× buffering the average DRAM bandwidth requirement is 6 bits per cycle but the peak bandwidth requirement is 12 bits per cycle. The amount of buffering used will depend on the DRAM bandwidth available to the DWU unit.

Should the DWU fail to get the required DRAM access within the specified time, the DWU will stall the DNC data generation. The DWU will issue the stall in sufficient time for the DNC to respond and still not cause a FIFO overrun. Should the stall persist for a sufficiently long time, the PHI will be starved of data and be unable to deliver data to the printhead in time. The sizing of the dotline store FIFO and internal FIFOs should be chosen so as to prevent such a stall happening.

30.7 Dotline Data in Memory

The dot data shift register order in the printhead is shown in FIG. 254 (the transmit order is the opposite of the shift register order). In the example the type 0 printhead IC transmit order is increasing even color data followed by decreasing odd color data. The type 1 printhead IC transmit order is decreasing odd color data followed by increasing even color data. For both printhead ICs the even data is always increasing order and odd data is always decreasing. The PHI controls which printhead IC data gets shifted to.

From this it is beneficial to store even data in increasing order in DRAM and odd data in decreasing order. While this order suits the example printhead, other printheads exist where it would be beneficial to store even data in decreasing order, and odd data in increasing order, hence the order is configurable. The order that data is stored in memory is controlled by setting the ColorLineSense register.

The dot order in DRAM for increasing and decreasing sense is shown in FIG. 260 and FIG. 261 respectively. For each line in the dot store the order is the same (although for odd lines the numbering will be different the order will remain the same). Dot data from the DNC is always received in increasing dot number order. For increasing sense dot data is bundled into 256-bit words and written in increasing order in DRAM, word 0 first, then word 1, and so on to word N, where N is the number of words in a line.

For decreasing sense dot data is also bundled into 256-bit words, but is written to DRAM in decreasing order, i.e. word N is written first then word N−1 and so on to word 0. For both increasing and decreasing sense the data is aligned to bit 0 of a word, i.e. increasing sense always starts at bit 0, decreasing sense always finishes at bit 0.

Each half color is configured independently of any other color. The ColorBaseAdr register specifies the position where data for a particular dotline FIFO will begin writing to. Note that for increasing sense colors the ColorBaseAdr register specifies the address of the first word of first line of the fifo, whereas for decreasing sense colors the ColorBaseAdr register specifies the address of last word of the first line of the FIFO.

Dot data received from the DNC is bundled in 256-bit words and transferred to the DRAM. Each line of data is stored consecutively in DRAM, with each line separated by ColorLineInc number of words.

For each line stored in DRAM the DWU increments the line count and calculates the DRAM address for the next line to store.

This process continues until ColorFifoSize number of lines are stored, after which the DRAM address will wrap back to the ColorBaseAdr address.

As each line is written to the FIFO, the DWU increments the FifoFillLevel register, and as the LLU reads a line from the FIFO the FifoFillLevel register is decremented. The LLU indicates that it has completed reading a line by a high pulse on the llu_dwu_line_rd line.

When the number of lines stored in the FIFO is equal to the MaxWriteAhead value the DWU will indicate to the DNC that it is no longer able to receive data (i.e. a stall) by deasserting the dwu_dnc_ready signal.

The ColorEnable register determines which color planes should be processed, if a plane is turned off, data is ignored for that plane and no DRAM accesses for that plane are generated.

30.8 Specifying Dot FIFOs

The dot line FIFOs when accessed by the LLU are specified differently than when accessed by the DWU. The DWU uses a start address and number of lines value to specify a dot FIFO, the LLU uses a start and end address for each dot FIFO. The mechanisms differ to allow more efficient implementations in each block.

As a result of limitations in the LLU the dot FIFOs must be specified contiguously and increasing in DRAM. See section 31.6 on page 504 for further information.

30.9 Implementation

30.9.1 Definitions of I/O

TABLE 206 DWU I/O Definition Port name Pins I/O Description Clocks and Resets Pclk 1 In System Clock prst_n 1 In System reset, synchronous active low DNC Interface dwu_dnc_ready 1 Out Indicates that DWU is ready to accept data from the DNC. dnc_dwu_avail 1 In Indicates valid data present on dnc_dwu_data. dnc_dwu_data[5:0] 6 In Input bi-level dot data in 6 ink planes. LLU Interface dwu_llu_line_wr 1 Out DWU line write. Indicates that the DWU has completed a full line write. Active high llfu_dwu_line_rd 1 In LLU line read. Indicates that the LLU has completed a line read. Active high. PCU Interface pcu_dwu_sel 1 In Block select from the PCU. When pcu_dwu_sel is high both pcu_adr and pcu_dataout are valid. pcu_rwn 1 In Common read/not-write signal from the PCU. pcu_adr[7:2] 6 In PCU address bus. Only 6 bits are required to decode the address space for this block. pcu_dataout[31:0] 32 In Shared write data bus from the PCU. dwu_pcu_rdy 1 Out Ready signal to the PCU. When dwu_pcu_rdy is high it indicates the last cycle of the access. For a write cycle this means pcu_dataout has been registered by the block and for a read cycle this means the data on dwu_pcu_datain is valid. dwu_pcu_datain[31:0] 32 Out Read data bus to the PCU. DIU Interface dwu_diu_wreq 1 Out DWU requests DRAM write. A write request must be accompanied by a valid write address together with valid write data and a write valid. dwu_diu_wadr[21:5] 17 Out Write address to DIU 17 bits wide (256-bit aligned word) diu_dwu_wack 1 In Acknowledge from DIU that write request has been accepted and new write address can be placed on dwu_diu_wadr dwu_diu_data[63:0] 64 Out Data from DWU to DIU. 256-bit word transfer over 4 cycles First 64-bits is bits 63:0 of 256 bit word Second 64-bits is bits 127:64 of 256 bit word Third 64-bits is bits 191:128 of 256 bit word Fourth 64-bits is bits 255:192 of 256 bit word dwu_diu_wvalid 1 Out Signal from DWU indicating that data on dwu_diu_data is valid. 30.9.2 DWU Partition 30.9.3 Configuration Registers

The configuration registers in the DWU are programmed via the PCU interface. Refer to section 21.8.2 on page 321 for a description of the protocol and timing diagrams for reading and writing registers in the DWU. Note that since addresses in SoPEC are byte aligned and the PCU only supports 32-bit register reads and writes, the lower 2 bits of the PCU address bus are not required to decode the address space for the DWU. When reading a register that is less than 32 bits wide zeros should be returned on the upper unused bit(s) of dwu_pcu_data. Table 207 lists the configuration registers in the DWU.

TABLE 207 DWU registers description Address DWU_base+ Register #bits Reset Description Control Registers 0x00 Reset 1 0x1 Active low synchronous reset, self de- activating. A write to this register will cause a DWU block reset. 0x04 Go 1 0x0 Active high bit indicating the DWU is programmed and ready to use. A low to high transition will cause DWU block internal states to reset (configuration registers are not reset). Dot Line Store Configuration 0x08-0x34 ColorBaseA  12 × 17 0x000 00 Specifies the base address (in words) in dr[11:0][21:5] memory where data from a particular half color (N) will be placed. For increasing sense colors the ColorBase- Adr register specifies the address of the first word of first line of the fifo, whereas for decreasing sense colors the ColorBaseAdr register specifies the address of last word of the first line of the fifo. 0x38-0x64 ColorFifoSize[11:0] 12 × 8 0x00 Indicates the number of lines in the FIFO before the line increment will wrap around in memory. Bus 0,1 - Even, Odd line color 0 Bus 2, 3 - Even, Odd line color 1 Bus 4, 5 - Even, Odd line color 2 Bus 6, 7 - Even, Odd line color 3 Bus 8, 9 - Even, Odd line color 4 Bus 10, 11 - Even, Odd line color 5 0x68 ColorLineSense 2 0x2 Specifies whether data written to DRAM for this half color is increasing or decreasing sense 0 - Decreasing sense 1 - Increasing sense Bit 0 Defines even color sense, Bit 1 Defines odd color sense. 0x6C ColorEnable 6 0x3F Indicates whether a particular color is active or not. When inactive no data is written to DRAM for that color. 0 - Color off 1 - Color on One bit per color, bit 0 is Color 0 and so on. 0x70 MaxWriteAhead 8 0x00 Specifies the maximum number of lines that the DWU can be ahead of the LLU 0x74 LineSize 16 0x000 0 Indicates the number of dots per line produced by the DWU. 0x78 MaxNozzleSkew 4 0x0 Specifies the number of dot-pairs the DWU needs to generate to flush the data skew buffers. Corresponds to the non-printable area of the printhead. 0x7C-0xA8 NozzleSkew 12 × 4 0x0 Specifies the relative skew of dot data nozzle rows in the printhead. Valid range is 0 (no skew) through to 12. Units represent dot-pairs, a skew of 1 for a row represents two dots on the page. Bus 0,1 - Even, Odd line color 0 Bus 2, 3 - Even, Odd line color 1 Bus 4, 5 - Even, Odd line color 2 Bus 6, 7 - Even, Odd line color 3 Bus 8, 9 - Even, Odd line color 4 Bus 10, 11 - Even, Odd line color 5 0xAC ColorLineInc 8 0x00 Specifies the number of words (256-bit words) per dot line - 1. Working Registers 0xB0 LineDotCnt 16 0x000 0 Indicates the number of remaining dots in the current line. (Read Only) 0xB4 FifoFillLevel 8 0x00 Number of lines in the FIFO, written to but not read. (Read Only)

A low to high transition of the Go register causes the internal states of the DWU to be reset. All configuration registers will remain the same. The block indicates the transition to other blocks via the dwu_go_pulse signal.

30.9.4 Data Skew

The data skew block inserts the shape of the printhead join into the dot data stream by delaying dot data by the relative nozzle skew amount (given by nozzle_skew). It generates zero fill data introduced introduced into the dot data stream to achieve the relative skew (and also to flush dot data from the delay registers).

The data skew block consists of 12 12-bit shift registers, one per color odd and even. The shift registers are in groups of 6, one group for even colors, and one for odd colors. Each time a valid data word is received from the DNC the dot data is shifted into either the odd or even group of shift registers. The odd_even_sel register determines which group of shift registers are valid for that cycle and alternates for each new valid data word. When a valid word is received for a group of shift registers, the shift register is shifted by one location with the new data word shifted into the registers (the top word in the register will be discarded).

When the dot counter determines that the data skew block should zero fill (zero_fill), the data skew block will shift zero dot data into the shift registers until the line has completed. During this time the DNC will be stalled by the de-assertion of the dwu_dnc_ready signal.

The data skew block selects dot data from the shift registers and is passed to the buffer address generator block. The data bits selected is determined by the configured index values in the NozzleSkew registers.

// determine when data is valid data_valid = (((dnc_dwu_avail == 1)OR(zero_fill == 1)) AND (dwu_ready ==1)) // implement the zero fill mux if (zero_fill == 1) then dot_data_in = 0 else dot_data_in = dnc_dwu_data // the data delay buffers if (dwu_go_pulse ==1) then data_delay[1:0][11:0][5:0] = 0 // reset all delay buffer odd=1,even=0 odd_even_sel = 0 elsif (data_valid == 1) then { odd_even_sel = ~odd_even_sel // update the odd/even buffers, with shift data_delay[odd_even_sel][11:1][5:0] = data_delay[odd_even_sel][10:0][5:0] // shift data data delay[odd even sel][0][5:0] = dot data in [5:0] // shift in new data // select the correct output data for (i=0;i<6; i++) { // skew selector skew = nozzle skew[ {i,odd even sel} ] // temporary variable // data select array, include data delay and input dot data data_select[12:0] = {data_delay[odd_even_sel][11:0], dot_data_in} // mux output the data word to next block (13 to 1 mux) dot_data[i] = data_select[skew][i] } } 30.9.5 Fifo Fill Level

The DWU keeps a running total of the number of lines in the dot store FIFO. Each time the DWU writes a line to DRAM (determined by the DIU interface subblock and signalled via line_wr) it increments the filllevel and signals the line increment to the LLU (pulse on dwu_llu_line_wr). Conversely if it receives an active llu_dwu_line_rd pulse from the LLU, the filllevel is decremented. If the filllevel increases to the programmed max level (max_write_ahead) then the DWU stalls and indicates back to the DNC by de-asserting the dwu_dnc_ready signal.

If one or more of the DIU buffers fill, the DIU interface signals the fill level logic via the buf_full signal which in turn causes the DWU to de-assert the dwu_dnc_ready signal to stall the DNC. The buf_full signals will remain active until the DIU services a pending request from the full buffer, reducing the buffer level.

When the dot counter block detects that it needs to insert zero fill dots (zero_fill equals 1) the DWU will stall the DNC while the zero dots are being generated (by de-asserting dwu_dnc_ready), but will allow the data skew block to generate zero fill data (the dwu_ready signal).

dwu_dnc_ready = ~((buf_full== 1) OR (filllevel == max_write_ahead ) OR (zero_fill == 1)) dwu_ready = ~((buf_full== 1) OR (filllevel == max_write_ahead ))

The DWU does not increment the fill level until a complete line of dot data is in DRAM not just a complete line received from the DNC. This ensures that the LLU cannot start reading a partial line from DRAM before the DWU has finished writing the line.

The fill level is reset to zero each time a new page is started, on receiving a pulse via the dwu_go_pulse signal.

The line fifo fill level can be read by the CPU via the PCU at any time by accessing the FifoFillLevel register.

30.9.6 Buffer Address Generator

30.9.6.1 Buffer Address Generator Description

The buffer address generator subblock is responsible for accepting data from the data skew block and writing it to the DIU buffers in the correct order.

The buffer address and active bit-write for a particular dot data write is calculated by the buffer address generator based on the dot count of the current line, programmed sense of the color and the line size.

All configuration registers should be programmed while the Go bit is set to zero, once complete the block can be enabled by setting the Go bit to one. The transition from zero to one will cause the internal states to reset.

If the color_line_sense signal for a color is one (i.e. increasing) then the bit-write generation is straight forward as dot data is aligned with a 256-bit boundary. So for the first dot in that color, the bit 0 of the wr_bit bus will be active (in buffer word 0), for the second dot bit 1 is active and so on to the 255^(th) dot where bit 63 is active (in buffer word 3). This is repeated for all 256-bit words until the final word where only a partial number of bits are written before the word is transferred to DRAM.

If color_line_sense signal for a color is zero (i.e. decreasing) the bit-write generation for that color is adjusted by an offset calculated from the pre-programmed line length (line_size). The offset adjusts the bit write to allow the line to finish on a 256-bit boundary. For example if the line length was 400, for the first dot received bit 7 (line length is halved because of odd/even lines of color) of the wr_bit is active (buffer word 3), the second bit 6 (buffer word 3), to the 200^(th) dot of data with bit 0 of wr_bit active (buffer word 0).

30.9.6.2 Bit-Write Decode

The buffer address generator contains 2 instances of the bit-write decode, one configured for odd dot data the other for even. The counter (either up or down counter) used to generate the addresses is selected by the color_line_sense signal. Each block determines if it is active on this cycle by comparing its configured type with the current dot count address and the data_active signal.

The wr_bit bus is a direct decoding of the lower 6 count bits (count[6:1]), and the DIU buffer address is the remaining higher bits of the counter (count[10:7]).

The signal generation is given as follows:

// determine the counter to use if (color_line_sense == 1 ) count = up_cnt[10:0] else count = dn_cnt[10:0] // determine if active, based on instance type wr_en = data_active & (count[0] {circumflex over ( )} odd_even_type) // odd =1, even =0 // determine the bit write value wr_bit[63:0] = decode(count[6:1]) // determine the buffer 64-bit address wr_adr[3:0] = count[10:7] 30.9.6.3 Up Counter Generator

The up counter increments for each new dot and is used to determine the write position of the dot in the DIU buffers for increasing sense data. At the end of each line of dot data (as indicated by line_fin), the counter is rounded up to the nearest 256-bit word boundary. This causes the DIU buffers to be flushed to DRAM including any partially filled 256-bit words. The counter is reset to zero if the dwu_go_pulse is one.

// Up-Counter Logic if (dwu_go_pulse == 1) then { up_cnt[10:0] = 0 elsif (line_fin == 1 ) then // round up if (up_cnt[8:1] != 0) up_cnt[10:9]++ else up_cnt[10:9] // bit-selector up_cnt[7:0]=0 elsif (data_valid == 1) then up_cnt[7:0]++ 30.9.6.4 Down Counter Generator

The down counter logic decrements for each new dot and is used to determine the write position of the dot in the DUI buffers for decreasing sense data. When the dwu_go_pulse bit is one the lower bits (i.e. 8 to 0) of the counter are reset to line size value (line_size), and the higher bits to zero. The bits used to determine the bit-write values and 64-bit word addresses in the DIU buffers begin at line size and count down to zero. The remaining higher bits are used to determine the DIU buffer 256-bit address and buffer fill level, begin at zero and count up. The counter is active when valid dot data is present, i.e. data_valid equals 1.

When the end of line is detected (line_fin equals 1) the counter is rounded to the next 256-bit word, and the lower bits are reset to the line size value.

//Down-Counter Logic if (dwu_go_pulse == 1) then dn_cnt[8:0] = line_size[8:0] dn_cnt[10:9] = 0 elsif (line_fin == 1 ) then // perform rounding up if (dn_cnt[8:1] != 0) dn_cnt[10:9]++ else dn_cnt[10:9] // bit-select is reset dn_cnt[8:0]=line_size[8:0] // bit select bits elsif (data_valid == 1) then dn_cnt[8:0] −− dn_cnt[10:9]++ 30.9.6.5 Dot Counter

The dot counter simply counts each active dot received from the data skew block. It sets the counter to line_size and decrements each time a valid dot is received. When the count equals zero the line_fin signal is pulsed and the counter is reset to line_size.

When the count is less than the max_nozzle_skew* 2 value the dot counter indicates to the data skew block to zero fill the remainder of the line (via the zero_fill signal). Note that the max_nozzle_skew units are dot-pairs as opposed to dots, hence the by 2 multiplication for comparison with the dot counter.

The counter is reset to line_size when dwu_go_pulse is 1.

30.9.7 DIU Buffer

The DIU buffer is a 64 bit×8 word dual port register array with bit write capability. The buffer could be implemented with flip-flops should it prove more efficient.

30.9.8 DIU Interface

30.9.8.1 DIU Interface General Description

The DIU interface determines when a buffer needs a data word to be transferred to DRAM. It generates the DRAM address based on the dot line position, the color base address and the other programmed parameters. A write request is made to DRAM and when acknowledged a 256-bit data word is transferred. The interface determines if further words need to be transferred and repeats the transfer process.

If the FIFO in DRAM has reached its maximum level, or one of the buffers has temporarily filled, the DWU will stall data generation from the DNC.

A similar process is repeated for each line until the end of page is reached. At the end of a page the CPU is required to reset the internal state of the block before the next page can be printed. A low to high transition of the Go register will cause the internal block reset, which causes all registers in the block to reset with the exception of the configuration registers. The transition is indicated to subblocks by a pulse on dwu_go_pulse signal.

30.9.8.2 Interface Controller

The interface controller state machine waits in Idle state until an active request is indicated by the read pointer (via the req_active signal). When an active request is received the machine proceeds to the ColorSelect state to determine which buffers need a data transfer. In the ColorSelect state it cycles through each color and determines if the color is enabled (and consequently the buffer needs servicing), if enabled it jumps to the Request state, otherwise the color_cnt is incremented and the next color is checked.

In the Request state the machine issues a write request to the DIU and waits in the Request state until the write request is acknowledged by the DIU (diu_dwu_wack). Once an acknowledge is received the state machine clocks through 4 cycles transferring 64-bit data words each cycle and incrementing the corresponding buffer read address. After transferring the data to the DIU the machine returns to the ColorSelect state to determine if further buffers need servicing. On the transition the controller indicates to the address generator (adr_update) to update the address for that selected color.

If all colors are transferred (color_cnt equal to 6) the state machine returns to Idle, updating the last word flags (group_fin) and request logic (req_update).

The dwu_diu_wvalid signal is a delayed version of the buf_rd_en signal to allow for pipeline delays between data leaving the buffer and being clocked through to the DIU block.

The state machine will return from any state to Idle if the reset or the dwu_go_pulse is 1.

30.9.8.3 Address Generator

The address generator block maintains 12 pointers (color adr[11:0]) to DRAM corresponding to current write address in the dot line store for each half color. When a DRAM transfer occurs the address pointer is used first and then updated for the next transfer for that color. The pointer used is selected by the req_sel bus, and the pointer update is initiated by the adr_update signal from the interface controller.

The pointer update is dependent on the sense of the color of that pointer, the pointer position in a line and the line position in the FIFO. The programming of the color_base_adr needs to be adjusted depending of the sense of the colors. For increasing sense colors the color_base_adr specifies the address of the first word of first line of the fifo, whereas for decreasing sense colors the color_base_adr specifies the address of last word of the first line of the FIFO.

For increasing colors, the initialization value (i.e. when dwu_go_pulse is 1) is the color_base_adr.

For each word that is written to DRAM the pointer is incremented. If the word is the last word in a line (as indicated by last_wd from that read pointers) the pointer is also incremented. If the word is the last word in a line, and the line is the last line in the FIFO (indicated by fifo_end from the line counter) the pointer is reset to color_base_adr.

In the case of decreasing sense colors, the initialization value (i.e. when dwu_go_pulse is 1) is the color_base_adr. For each line of decreasing sense color data the pointer starts at the line end and decrements to the line start. For each word that is written to DRAM the pointer is decremented. If the word is the last word in a line the pointer is incremented by color_line_inc*2+1. One line length to account for the line of data just written, and another line length for the next line to be written. If the word is the last word in a line, and the line is the last line in the FIFO the pointer is reset to the initialization value (i.e. color_base_adr).

The address is calculated as follows:

if (dwu_go_pulse == 1) then color_adr[11:0] = color_base_adr[11:0][21:5] elsif (adr_update == 1) then { // determine the color color = req_sel[3:0] // line end and fifo wrap if ((fifo_end[color] == 1) AND (last_wd == 1)) then { // line end and fifo wrap color_adr[color] = color_base_adr[color][21:5] } elsif ( last_wd == 1) then { // just a line end no fifo wrap if (color_line_sense[color % 2] == 1) then // increasing sense color_adr[color] ++ else // decreasing sense color_adr[color] = color_adr[color] + ( color_line_inc * 2) + 1 } else { // regular word write if (color_line_sense[color % 2] == 1) then // increasing sense color_adr[color]++ else // decreasing sense color_adr[color]−− } } // select the correct address, for this transfer dwu_diu_wadr = color_adr[req_sel] 30.9.8.4 Line Count

The line counter logic counts the number of dot data lines stored in DRAM for each color. A separate pointer is maintained for each color. A line pointer is updated each time the final word of a line is transferred to DRAM. This is determined by a combination of adr_update and last_wd signals. The pointer to update is indicated by the req_sel bus.

When an update occurs to a pointer it is compared to zero, if it is non-zero the count is decremented, otherwise the counter is reset to color_fifo_size. If a counter is zero the fifo_end signals is set high to indicates to the address generator block that the line is the last line of this colors fifo.

If the dwu_go_pulse signal is one the counters are reset to color_fifo_size.

if (dwu_go_pulse == 1) then line_cnt[11:0] = color_fifo_size[11:0] elsif ((adr_update == 1) AND (last_wd == 1)) then { // determine the pointer to operate on color = req_sel[3:0] // update the pointer if (line_cnt[color] == 0) then line_cnt[color] = color_fifo_size[color] else line_cnt[i] −− } // count is zero its the last line of fifo for(i=0 ;i <12;i++){ fifo_end[i] = (line_cnt[i] == 0) } 30.9.8.5 Read Pointer

The read pointer logic maintains the buffer read address pointers. The read pointer is used to determine which 64-bit words to read from the buffer for transfer to DRAM.

The read pointer logic compares the read and write pointers of each DIU buffer to determine which buffers require data to be transferred to DRAM, and which buffers are full (the buf_full signal).

Buffers are grouped into odd and even buffers groups. If an odd buffer requires DRAM access the odd_pend signals will be active, if an even buffer requires DRAM access the even_pend signals will be active. If both odd and even buffers require DRAM access at exactly the same time, the even buffers will get serviced first. If a group of odd buffers are being serviced and an even buffer becomes pending, the odd group of buffers will be completed before the starting the even group, and vice versa.

If any buffer requires a DRAM transfer, the logic will indicate to the interface controller via the req_active signal, with the odd_even_sel signal determining which group of buffers get serviced. The interface controller will check the color_enable signal and issue DRAM transfers for all enabled colors in a group. When the transfers are complete it tells the read pointer logic to update the requests pending via req update signal.

The req_sel[3:0] signal tells the address generator which buffer is being serviced, it is constructed from the odd_even_sel signal and the color_cnt[2:0] bus from the interface controller. When data is being transferred to DRAM the word pointer and read pointer for the corresponding buffer are updated. The req_sel determines which pointer should be incremented.

// determine if request is active even if ( wr_adr[0][3:2] != rd_adr[0][3:2] ) even_pend = 1 else even_pend = 0 // determine if request is active odd if ( wr_adr[1][3:2] != rd_adr[1][3:2] ) even_pend = 1 else even_pend = 0 // determine if any buffer is full if ((wr_adr[0][3:0] − rd_adr[0][3:0]) > 7)OR((wr_adr[1][3:0] − rd_adr[1][3:0])> 7)) then buf_full = 1 // fixed servicing order, only update when controller dictates so if (req_update == 1) then { if (even_pend == 1) then // even always first odd_even_sel = 0 req_active = 1 elsif (odd_pend == 1 ) then // then check odd odd_even_sel = 0 req_active = 1 else // nothing active odd_even_sel = 0 req_active = 0 } // selected requestor req_sel[3:0] = {color_cnt[2:0] , odd_even_sel} // concatentation

The read address pointer logic consists of 2 2-bit counters and a word select pointer. The pointers are reset when dwu_go_pulse is one. The word pointer (word_ptr) is common to all buffers and is used to read out the 64-bit words from the DIU buffer. It is incremented when buf_rd_en is active. When a group of buffers are updated the state machine increments the read pointer (rd_ptr[odd_even_sel]) via the group_fin signal. A concatenation of the read pointer and the word pointer are use to construct the buffer read address. The read pointers are not reset at the end of each line.

// determine which pointer to update if (dwu_go_pulse == 1) then rd_ptr[1:0] = 0 word_ptr = 0 elsif (buf_rd_en == 1) then { word_ptr++ // word pointer update elsif (group_fin == 1) then rd_ptr [odd_even_sel]++ // update the read pointer // create the address from the pointer, and word reader rd_adr[odd_even_sel] = {rd_ptr[odd_even_sel],word_ptr} // concatenation

The read pointer block determines if the word being read from the DIU buffers is the last word of a line. The buffer address generator indicate the last dot is being written into the buffers via the line fin signal. When received the logic marks the 256-bit word in the buffers as the last word. When the last word is read from the DIU buffer and transferred to DRAM, the flag for that word is reflected to the address generator.

// line end set the flags if (dwu_go_pulse == 1) then last_flag[1:0][1:0] = 0 elsif (line_fin == 1 ) then // determines the current 256-bit word even been written to last_flag[0][wr_adr[0][2]] = 1 // even group flag // determines the current 256-bit word odd been written to last_flag[1][wr_adr[1][2]] = 1 // odd group flag // last word reflection to address generator last_wd = last_flag[odd_even_sel][rd_ptr[req_sel][0]] // clear the flag if (group_fin == 1 ) then last_flag[odd_even_sel][rd_ptr[req_sel][0]] = 0

When a complete line has been written into the DIU buffers (but has not yet been transferred to DRAM), the buffer address generator block will pulse the line_fin signal. The DWU must wait until all enabled buffers are transferred to DRAM before signaling the LLU that a complete line is available in the dot line store (dwu_llu_line_wr signal). When the line_fin is received all buffers will require transfer to DRAM. Due to the arbitration, the even group will get serviced first then the odd. As a result the line finish pulse to the LLU is generated from the last_flag of the odd group.

// must be odd,odd group transfer complete and the last word dwu_llu_line_wr = odd_even_sel AND group_fin AND last_wd 31 Line Loader Unit (LLU) 31.1 Overview

The Line Loader Unit (LLU) reads dot data from the line buffers in DRAM and structures the data into even and odd dot channels destined for the same print time. The blocks of dot data are transferred to the PHI and then to the printhead. FIG. 267 shows a high level data flow diagram of the LLU in context.

31.2 Physical Requirement Imposed by the Printhead

The DWU re-orders dot data into 12 separate dot data line FIFOs in the DRAM. Each FIFO corresponds to 6 colors of odd and even data. The LLU reads the dot data line FIFOs and sends the data to the printhead interface. The LLU decides when data should be read from the dot data line FIFOs to correspond with the time that the particular nozzle on the printhead is passing the current line. The interaction of the DWU and LLU with the dot line FIFOs compensates for the physical spread of nozzles firing over several lines at once. For further explanation see Section 30 Dotline Writer Unit (DWU) and Section 32 Printhead Interface (PHI). FIG. 268 shows the physical relationship of nozzle rows and the line time the LLU starts reading from the dot line store.

Within each line of dot data the LLU is required to generate an even and odd dot data stream to the PHI block. FIG. 269 shows the even and dot streams as they would map to an example bi-lithic printhead. The PHI block determines which stream should be directed to which printhead IC.

31.3 Dot Generate and Transmit Order

The structure of the printhead ICs dictate the dot transmit order to each printhead IC. The LLU reads data from the dot line FIFO, generates an even and odd dot stream which is then re-ordered (in the PHI) into the transmit order for transfer to the printhead.

The DWU separates dot data into even and odd half lines for each color and stores them in DRAM. It can store odd or even dot data in increasing or decreasing order in DRAM. The order is programmable but for descriptive purposes assume even in increasing order and odd in decreasing order. The dot order structure in DRAM is shown in FIG. 261.

The LLU contains 2 dot generator units. Each dot generator reads dot data from DRAM and generates a stream of odd or even dots. The dot order may be increasing or decreasing depending on how the DWU was programmed to write data to DRAM. An example of the even and odd dot data streams to DRAM is shown in FIG. 270. In the example the odd dot generator is configured to produce odd dot data in decreasing order and the even dot generator produces dot data_in increasing order.

The PHI block accepts the even and odd dot data streams and reconstructs the streams into transmit order to the printhead.

The LLU line size refers to the page width in dots and not necessarily the printhead width. The page width is often the dot margin number of dots less than the printhead width. They can be the same size for full bleed printing.

31.4 LLU Start-Up

At the start of a page the LLU must wait for the dot line store in DRAM to fill to a configured level (given by FifoReadThreshold) before starting to read dot data. Once the LLU starts processing dot data for a page it must continue until the end of a page, the DWU (and other PEP blocks in the pipeline) must ensure there is always data in the dot line store for the LLU to read, otherwise the LLU will stall, causing the PHI to stall and potentially generate a print error. The FifoReadThreshold should be chosen to allow for data rate mismatches between the DWU write side and the LLU read side of the dot line FIFO. The LLU will not generate any dot data until FifoReadThreshold level in the dot line FIFO is reached.

Once the FifoReadThreshold is reached the LLU begins page processing, the FifoReadThreshold is ignored from then on.

When the LLU begins page processing it produces dot data for all colors (although some dot data color may be null data). The LLU compares the line count of the current page, when the line count exceeds the ColorRelLine configured value for a particular color the LLU will start reading from that colors FIFO in DRAM. For colors that have not exceeded the ColorRelLine value the LLU will generate null data (zero data) and not read from DRAM for that color. ColorRelLine[N] specifies the number of lines separating the N^(th) half color and the first half color to print on that page. For the example printhead shown in FIG. 268, color 0 odd will start at line 0, the remaining colors will all have null data. Color 0 odd will continue with real data until line 5, when color 0 odd and even will contain real data the remaining colors will contain null data. At line 10, color 0 odd and even and color 1 odd will contain real data, with remaining colors containing null data. Every 5 lines a new half color will contain real data and the remaining half colors null data until line 55, when all colors will contain real data. In the example ColorRelLine[O]=5, ColorRelLine[1]=0, ColorRelLine[2]=15, ColorRelLine[3]=10. etc.

It is possible to turn off any one of the color planes of data (via the ColorEnable register), in such cases the LLU will generate zeroed dot data information to the PHI as normal but will not read data from the DRAM.

31.4.1 LLU Bandwidth Requirements

The LLU is required to generate data for feeding to the printhead interface, the rate required is dependent on the printhead construction and on the line rate configured. The maximum data rate the LLU can produce is 12 bits of dot data per cycle, but the PHI consumes at 12 bits every 2 pclk cycles out of 3, i.e. 8 bits per pclk cycle. Therefore the DRAM bandwidth requirement for a double buffered LLU is 8 bits per cycle on average. If 1.5 buffering is used then the peak bandwidth requirement is doubled to 16 bits per cycle but the average remains at 8 bits per cycle. Note that while the LLU and PHI could produce data at the 8 bits per cycle rate, the DWU can only produce data at 6 bits per cycle rate.

31.5 Vertical Row Skew

Due to construction limitations of the bi-lithic printhead it is possible that nozzle rows may be misaligned relative to each other. Odd and even rows, and adjacent color rows may be horizontally misaligned by up to 2 dot positions. Vertical misalignment can also occur between both printhead ICs used to construct the printhead. The DWU compensates for the horizontal misalignment (see Section 30.5), and the LLU compensates for the vertical misalignment. For each color odd and even the LLU maintains 2 pointers into DRAM, one for feeding printhead A (CurrentPtrA) and other for feeding printhead B (CurrentPtrB). Both pointers are updated and incremented in exactly the same way, but differ in their initial value programming. They differ by vertical skew number of lines, but point to the same relative position within a line.

At the start of a line the LLU reads from the FIFO using CurrentPtrA until the join point between the printhead ICs is reached (specified by JoinPoint), after which the LLU reads from DRAM using CurrentPtrB. If the JoinPoint coincides with a 256-bit word boundary, the swap over from pointer A to pointer B is straightforward. If the JoinPoint is not on a 256-bit word boundary, the LLU must read the 256-bit word of data from CurrentPtrA location, generate the dot data up to the join point and then read the 256-bit word of data from CurrentPtrB location and generate dot data from the join point to the word end. This means that if the JoinPoint is not on a 256-bit boundary then the LLU is required to perform an extra read from DRAM at the join point and not increment the address pointers.

31.5.1 Dot Line FIFO Initialization

For each dot line FIFO there are 2 pointers reading from it, each skewed by a number of dot lines in relation to the other (the skew amount could be positive or negative). Determining the exact number of valid lines in the dot line store is complicated by two pointers reading from different positions in the FIFO. It is convenient to remove the problem by pre-zeroing the dot line FIFOs effectively removing the need to determine exact data validity. The dot FIFOs can be initialized in a number of ways, including

-   -   the CPU writing 0s,     -   the LBD/SFU writing a set of 0 lines (16 bits per cycle),     -   the HCU/DNC/DWU being programmed to produce 0 data         31.6 Specifying Dot FIFOS

The dot line FIFOs when accessed by the LLU are specified differently than when accessed by the DWU. The DWU uses a start address and number of lines value to specify a dot FIFO, the LLU uses a start and end address for each dot FIFO. The mechanisms differ to allow more efficient implementations in each block.

The start address for each half color N is specified by the ColorBaseAdr[N] registers and the end address (actually the end address plus 1) is specified by the ColorBaseAdr[N+1]. Note there are 12 colors in total, 0 to 11, the ColorBaseAdr[12] register specifies the end of the color 11 dot FIFO and not the start of a new dot FIFO. As a result the dot FIFOs must be specified contiguously and increasing in DRAM.

31.7 Implementation

31.7.1 LLU Partition

31.7.2 Definitions of I/O

TABLE 208 LLU I/O definition Port name Pins I/O Description Clocks and Resets Pclk 1 In System clock prst_n 1 In System reset, synchronous active low PHI Interface llu_phi_data[1:0][5:0] 2 × 6 Out Dot Data from LLU to the PHI, each bit is a color plane 5 down to 0. Bus 0 - Even dot data stream Bus 1 - Odd dot data stream Data is active when corresponding bit is active in llu_phi_avail bus phi_llu_ready[1:0] 2 In Indicates that PHI is ready to accept data from the LLU 0 - Even dot data stream 1 - Odd dot data stream llu_phi_avail[1:0] 2 Out Indicates valid data present on corresponding llu_phi_data. 0 - Even dot data stream 1 - Odd dot data stream DIU Interface llu_diu_rreq 1 Out LLU requests DRAM read. A read request must be accompanied by a valid read address. llu_diu_radr[21:5] 17 Out Read address to DIU 17 bits wide (256-bit aligned word). diu_llu_rack 1 In Acknowledge from DIU that read request has been accepted and new read address can be placed on llu_diu_radr diu_data[63:0] 64 In Data from DIU to LLU. Each access is 256-bits received over 4 clock cycles First 64-bits is bits 63:0 of 256 bit word Second 64-bits is bits 127:64 of 256 bit word Third 64-bits is bits 191:128 of 256 bit word Fourth 64-bits is bits 255:192 of 256 bit word diu_llu_rvalid 1 In Signal from DIU telling LLU that valid read data is on the diu_data bus DWU Interface dwu_llu_line_wr 1 In DWU line write. Indicates that the DWU has completed a full line write. Active high llu_dwu_line_rd 1 Out LLU line read. Indicates that the LLU has completed a line read. Active high. PCU Interface pcu_llu_sel 1 In Block select from the PCU. When pcu_llu_sel is high both pcu_adr and pcu_dataout are valid. pcu_rwn 1 In Common read/not-write signal from the PCU. pcu_adr[7:2] 6 In PCU address bus. Only 6 bits are required to decode the address space for this block. pcu_dataout[31:0] 32 In Shared write data bus from the PCU. llu_pcu_rdy 1 Out Ready signal to the PCU. When llu_pcu_rdy is high it indicates the last cycle of the access. For a write cycle this means pcu_dataout has been registered by the block and for a read cycle this means the data on llu_pcu_datain is valid. llu_pcu_datain[31:0] 32 Out Read data bus to the PCU. 31.7.3 Configuration Registers

The configuration registers in the LLU are programmed via the PCU interface. Refer to section 21.8.2 on page 321 for a description of the protocol and timing diagrams for reading and writing registers in the LLU. Note that since addresses in SoPEC are byte aligned and the PCU only supports 32-bit register reads and writes, the lower 2 bits of the PCU address bus are not required to decode the address space for the LLU. When reading a register that is less than 32 bits wide zeros should be returned on the upper unused bit(s) of llu_pcu_datain. Table 209 lists the configuration registers in the LLU.

TABLE 209 LLU registers description Address LLU_base+ Register #bits Reset Description Control Registers 0x00 Reset 1 0x1 Active low synchronous reset, self de- activating. A write to this register will cause a LLU block reset. 0x04 Go 1 0x0 Active high bit indicating the LLU is programmed and ready to use. A low to high transition will cause LLU block internal states to reset. Configuration 0x08-0x38 ColorBaseAdr[12:0][21:5] 13 × 17 0x000 00 Specifies the base address (in words) in memory where data from a particular half color (N) will be placed. Also specifies the end address + 1 (256- bit words) in memory where fifo data for a particular half color ends. For color N the start address is ColorBaseAdr[N] and the end address +1 is ColorBase- Adr[N+1] 0x3C ColorEnable 6 0x3F Indicates whether a particular color is active or not. When inactive no data is written to DRAM for that color. 0 - Color off 1 - Color on One bit per color, bit 0 is Color 0 and so on. 0x40 LineSize 16 0x000 0 Indicates the number of dots per line. 0x44 FifoReadThreshold 8 0x00 Specifies the number of lines that should be in the FIFO before the LLU starts reading. 0x48-0x74 ColorRelLine[11:0] 12 × 8  0x00 Specifies the relative number of lines to wait from the first before starting to read dot data from the corresponding dot data FIFO Bus 0,1 - Even, Odd line color 0 Bus 2, 3 - Even, Odd line color 1 Bus 4, 5 - Even, Odd line color 2 Bus 6, 7 - Even, Odd line color 3 Bus 8, 9 - Even, Odd line color 4 Bus 10, 11 - Even, Odd line color 5 0x78-0x7C JoinPoint  2 × 16 0x000 0 Specifies the join point in dots between both printhead ICs. Bus 0 - Even dot generator join point Bus 1 - Odd dot generator join point 0x80-0x84 JoinWord 2 × 8 0x00 Specifies the join point in words between both printhead ICs. Bus 0 - Even dot generator join point Bus 1 - Odd dot generator join point 0x90-0xBC CurrentAdrA[11:0][21:5] 12 × 17 0x000 0 Current Address pointers associated with printhead A Bus 0,1 - Even, Odd line color 0 Bus 2, 3 - Even, Odd line color 1 Bus 4, 5 - Even, Odd line color 2 Bus 6, 7 - Even, Odd line color 3 Bus 8, 9 - Even, Odd line color 4 Bus 10, 11 - Even, Odd line color 5 Working registers 0xC0-0xEC CurrentAdrB[11:0][21:5] 12 × 17 0x000 0 Current Address pointers associated with printhead B Bus 0,1 - Even, Odd line color 0 Bus 2, 3 - Even, Odd line color 1 Bus 4, 5 - Even, Odd line color 2 Bus 6, 7 - Even, Odd line color 3 Bus 8, 9 - Even, Odd line color 4 Bus 10, 11 - Even, Odd line color 5 Working registers Working Registers 0xF0 FifoFillLevel 8 0x00 Number of lines in the dot line FIFO, line written in but not read out. (Read Only)

A low to high transition of the Go register causes the internal states of the LLU to be reset. All configuration registers will remain the same. The block indicates the transition to other blocks via the llu_go_pulse signal.

31.7.4 Dot Generator

The dot generator block is responsible for reading dot data from the DIU buffers and sending the dot data in the correct order to the PHI block. The dot generator waits for llu_en signal from the fifo fill level block, once active it starts reading data from the 6 DIU buffers and generating dot data for feeding to the PHI.

In the LLU there are two instances of the dot generator, one generating odd data and the other generating even data.

At any time the ready bit from the PHI could be de-asserted, if this happens the dot generator will stop generating data, and wait for the ready bit to be re-asserted.

31.7.4.1 Dot Count

In normal operation the dot counter will wait for the llu_en and the ready to be active before starting to count. The dot count will produce data as long as the phi_llu_ready is active. If the phi_llu_ready signal goes low the count will be stalled.

The dot counter increments for each dot that is processed per line. It is used to determine the line finish position, and the bit select value for reading from the DIU buffers. The counter is reset after each line is processed (line_fin signal). It determines when a line is finished by comparing the dot count with the configured line size divided by 2 (note that odd numbers of dots will be rounded down).

// define the line finish if (dot_cnt [14:0] == line_size[15:1] )then line_fin = 1 else line_fin = 0 // determine if word is valid dot_active = ((llu_en == 1) AND (phi_llu_ready == 1) AND (buf_emp == 0)) // counter logic if (llu_go_pulse == 1) then dot_cnt = 0 elsif ((dot_active == 1)AND (line_fin == 1)) then dot_cnt = 0 elsif (dot_active == 1) then dot_cnt = dot_cnt + 1 else dot_cnt = dot_cnt // calculate the word select bits bit_sel[5:0] := dot_cnt[5:0]

The dot generator also maintains a read buffer pointer which is incremented each time a 64-bit word is processed. The pointer is used to address the correct 64-bit dot data word within the DIU buffers. The pointer is reset when llu_go_pulse is 1. Unlike the dot counter the read pointer is not reset each line but rounded up the nearest 256-bit word. This allows for more efficient use of the DIU buffers at line finish.

When the dot counter reaches the join point for the dot generator (join_point), it jumps to the next 256 bit word in the DIU buffer but continues to read from the next bit position within that word. If the join point coincides with a word boundary, no 256-bit increment is required.

// read pointer logic if (llu_go_pulse == 1) then read_adr = 0 elsif ((dot_active == 1)AND((dot_cnt[7:0] == 255)OR(line_fin == 1)))then // end of line round up read_adr[3:2] ++ read_adr[1:0] = 0 elsif ((dot_active == 1)AND(dot_cnt == join_point)AND(dot_cnt[5:0] == 63)) then // join point jump 256 bits read_adr[1:0] ++ // regular increment read_adr[3:2] ++ // join point 256 increment elsif ((dot_active == 1)AND(dot_cnt == join_point)AND(dot_cnt[5:0] != 63)) then // join point jump 256 bits, bottom bits remain the same read_adr[3:2] ++ // join point 256 increment only elsif ((dot_active == 1)AND(dot_cnt[5:0] == 63)) then read_adr [3:0] ++ // regular increment 31.7.5 Fifo Fill Level

The LLU keeps a running total of the number of lines in the dot line store FIFO. Every time the DWU signals a line end (dwu_llu_line_wr active pulse) it increments the filllevel. Conversely if the LLU detects a line end (line_rd pulse) the filllevel is decremented and the line read is signalled to the DWU via the llu_dwu_line_rd signal.

The LLU fill level block is used to determine when the dot line has enough data stored before the LLU should begin to start reading. The LLU at page start is disabled. It waits for the DWU to write lines to the dot line FIFO, and for the fill level to increase. The LLU remains disabled until the fill level has reached the programmed threshold (fifo read thres). When the threshold is reached it signals the LLU to start processing the page by setting llu_en high. Once the LLU has started processing dot data for a page it will not stop if the filllevel falls below the threshold, but will stall is filllevel falls to zero.

The line fifo filllevel can be read by the CPU via the PCU at any time by accessing the FifoFillLevel register. The CPU must toggle the Go register in the LLU for the block to be correctly initialized at page start and the fifo level reset to zero.

if (llu_go_pulse == 1) then filllevel = 0 elsif ((line_rd == 1) AND (dwu_llu_line_wr == 1)) then // do nothing elsif (line_rd == 1) then filllevel −− elsif (dwu_llu_line_wr == 1) then filllevel ++ // determine the threshold, and set the LLU going if (llu_go_pulse == 1) OR (filllevel == 0 )) then llu_en = 0 elsif (filllevel == fifo_read_threshold ) then llu_en = 1 31.7.6 DIU Interface 31.7.6.1 DIU Interface Description

The DIU interface block is responsible for determining when dot data needs to be read from DRAM, keeping the dot generators supplied with data and calculating the DRAM read address based on configured parameters, FIFO fill levels and position in a line.

The fill level block enables DIU requests by activating llu_en signal. The DIU interface controller then issues requests to the DIU for the LLU buffers to be filled with dot line data (or fill the LLU buffers with null data without requesting DRAM access, if required).

At page start the DIU interface determines which buffers should be filled with null data and which should request DRAM access. New requests are issued until the dot line is completely read from DRAM.

For each request to the DRAM the address generator calculates where in the DRAM the dot data should be read from. The color_enable bus determines which colors are enabled, the interface never issues DRAM requests for disabled colors.

31.7.6.2 Interface Controller

The interface controller co-ordinates and issues requests for data transfers from DRAM. The state machine waits in Idle state until it is enabled by the LLU controller (llu_en) and a request for data transfer is received from the write pointer block.

When an active request is received (req_active equals 1) the state machine jumps to the ColorSelect state to determine which colors (color_cnt) in the group need a data transfer. A group is defined as all odd colors or all even colors. If the color isn't enabled (color_enable) the count just increments, and no data is transferred. If the color is enabled, the state machine takes one of two options, either a null data transfer or an actual data transfer from DRAM. A null data transfer writes zero data to the DIU buffer and does not issue a request to DRAM.

The state machine determines if a null transfer is required by checking the color_start signal for that color.

If a null transfer is required the state machine doesn't need to issue a request to the DIU and so jumps directly to the data transfer states (Data0 to Data3). The machine clocks through the 4 states each time writing a null 64-bit data word to the buffer. Once complete the state machine returns to the ColorSelect state to determine if further transfers are required.

If the color_start is active then a data transfer is required. The state machine jumps to the Request state and issue a request to the DIU controller for DRAM access by setting llu_diu_rreq high. The DIU responds by acknowledging the request (diu_llu_rack equals 1) and then sending 4 64-bit words of data. The transition from Request to Data0 state signals the address generator to update the address pointer (adr_update). The state machine clocks through Data0 to Data3 states each time writing the 64-bit data into the buffer selected by the req_sel bus. Once complete the state machine returns to the ColorSelect state to determine if further transfers are required.

When in the ColorSelect state and all data transfers for colors in that group have been serviced (i.e. when color_cnt is 6) the state machine will return to the Idle state. On transition it will update the word counter logic (word_dec) and enabled the request logic (req_update).

A reset or llu_go_pulse set to 1 will cause the state machine to jump directly to Idle. The controller will remain in Idle state until it is enabled by the LLU controller via the llu_en signal. This prevents the DIU attempting the fill the DIU buffers before the dot line store FIFO has filled over its threshold level. μ31.7.6.3 Color Activate

The color activate logic maintains an absolute line count indicating the line number currently being processed by the LLU. The counter is reset when the llu_go_pulse is 1 and incremented each time a line_rd pulse is received. The count value (line_cnt) is used to determine when to start reading data for a color.

The count is implemented as follows:

if ( llu_go_pulse == 1) then line_cnt = 0 elsif ( line_rd == 1) then line_cnt ++

The color activate logic compares line count with the relative line value to determine when the LLU should start reading data from DRAM for a particular half color. It signals the interface controller block which colors are active for this dot line in a page (via the color_start bus). It is used by the interface controller to determine which DIU buffers require null data.

Once the color_start bit for a color is set it cannot be cleared in the normal page processing process. The bits must be reset by the CPU at the end of a page by transitioning the Go bit and causing a pulse on the llu_go_pulse signal.

Any color not enabled by the color_enable bus will never have its color_start bit set.

for (i=0; i<12;i++){ if ( llu_go_pulse == 1) then col_on[i] = 0 elsif ( color_enable[i % 6] == 1 ) then col_on[i] = 0 elsif ( line_cnt == color_rel_line[i]) then col_on[i] = 1 } // select either odd or even colors if ( odd_even_sel == 1 ) then // odd selected color_start[5:0] = {col_on[11],col_on[9],col_on[7],col_on[5],col_on[3],col_on[1 ]} else // even selected color_start[5:0] = {col_on[10],col_on[8],col_on[6],col_on[4],col_on[2],col_on[0 ]} 31.7.6.4 Address Generator

The address generator block maintains 24 pointers (current_adr_a[11:0] and current_adr_b[11:0]) to DRAM corresponding to 2 read addresses in the dot line FIFO for each half color. The current_adr a group of pointers are used when the dot generator is feeding printhead channel A, and the current_adr_b group of pointers are used when the dot generator is feeding printhead channel B. For each DRAM access the 2 address pointers are updated but only one can be used for an access. The word counter block determines which pointer group should be used to access DRAM, via the pointer select signals (ptr_sel). In certain cases (e.g. the join point is not 256-bit aligned and the word is on the join point) the address pointers should not be updated for an access, the word counter block determines the exception cases and indicates to the address generator to skip the update via the join_stall signal.

When a DRAM transfer occurs the address pointer is used first and then updated for the next transfer for the color. The pointer used is selected by the req_sel and ptr_sel buses, and the pointer update is initiated by the adr_update signal from the interface controller.

The address update is calculated as follows (pointer group A logic is shown but the same logic is used to update the B pointer group a clock cycle later):

// update the A pointers if (ptra_wr_en == 1) then // write from the configuration block current_adr_a[ptr_adr] = ptr_wr_data; elsif ( adr_update_a == 1) then { // address update from state machine if ((req_sel == NULL )OR (join_stall == 1)) then // do nothing else // temporary variable setup next_adr = current_adr_a[req_sel] + 1 start_adr = color_base_adr[req_sel] end_adr = color_base_adr[req_sel + 1] // determine how to update the pointer if (next_adr == end_adr) then current_adr_a[req_sel] = start_adr else current_adr_a[req_sel] = next_adr }

The correct address to use for a transfer is selected by the ptr_sel signals from the word counter block. They indicate which set of address pointers should be used based on the current word being transferred from the DRAM and the configured join point values (join_word).

// select the address pointer to use for access if (req_sel[0] == 1) then  // odd pointer selector if (ptr_sel[1] == 1) then llu_diu_radr = current_adr_b[req_sel] // latter part of line else llu_diu_radr = current_adr_a[req_sel] // former part of line else // even pointer selector if (ptr_sel[0] == 1) then llu_diu_radr = current_adr_b[req_sel] // latter part of line else llu_diu_radr = current_adr_a[req_sel] // former part of line 31.7.6.5 Write Pointer

The write pointer logic maintains the buffer write address pointers, determines when the DIU buffers need a data transfer and signals when the DIU buffers are empty. The write pointer determines the address in the DIU buffer that the data should be transferred to.

The write pointer logic compares the read and write pointers of each DIU buffer to determine which buffers require data to be transferred from DRAM, and which buffers are empty (the buf_emp signals).

Buffers are grouped into odd and even buffers, if an odd buffer requires DRAM access the odd_pend signals will be active, if an even buffer requires DRAM access the even_pend signals will be active. If both odd and even buffers require DRAM access at exactly the same time, the even buffers will get serviced first. If a group of odd buffers are being serviced and an even buffer becomes pending, the odd group of buffers will be completed before the starting the even group, and vice versa.

If any buffer requires a DRAM transfer, the logic will indicate to the interface controller via the req_active signal, with the odd_even_sel signal determining which group of buffers get serviced. The interface controller will check the color_enable signal and issue DRAM transfers for all enabled colors in a group. When the transfers are complete it tells the write pointer logic to update the request pending via req_update signal.

The req_sel[3:0] signal tells the address generator which buffer is being serviced, it is constructed from the odd_even_sel signal and the color_cnt[2:0] bus from the interface controller. When data is being transferred to DRAM the word pointer and write pointer for the corresponding buffer are updated. The req_sel determines which pointer should be incremented.

The write pointer logic operates the same way regardless of whether the transfer is null or not.

// determine which buffers need updates buf_emp[1:0] = 0 odd_pend = 0 even_pend = 0 if ( wr_adr[0][3:2] == rd_adr[0][3:2] ) even_pend = 1 if ( wr_adr[1][3:2] == rd_adr[1][3:2] ) odd_pend = 1 // determine if buffers are empty if ((wr_adr[0][3:0] == rd_adr[0][3:0])) then buf_emp[0] = 1 if ((wr_adr[1][3:0] == rd_adr[1][3:0])) then buf_emp[1] = 1 // fixed servicing order, only update when controller dictates so if (req_update == 1) then { if (even_pend == 1) then // even always first odd_even_sel = 0 req_active = 1 elsif (odd_pend == 1 ) then // then check odd odd_even_sel = 0 req_active = 1 else // nothing active odd_even_sel = 0 req_active = 0 } // selected requestor req_sel[3:0] = {color_cnt[2:0],odd_even_sel} // concatentation

The write address pointer logic consists of 2 2-bit counters and a word select pointer. The counters are reset when llu_go_pulse is one. The word pointer (word_ptr) is common to all buffers and is used to write 64-bit words into the DIU buffer. It is incremented when buf_rd_en is active.

When a group of buffers are updated the state machine increments the write pointer (wr_ptr[odd_even_sel]) via the group_fin signal. A concatenation of the write pointer and the word pointer are use to construct the buffer write address. The write pointers are not reset at the end of each line.

// determine which pointer to update if (llu_go_pulse == 1) then wr_ptr[1:0] = 0 word_ptr = 0 elsif (buf_rd_en == 1) then word_ptr++ wr_en[req_sel] = 1 elsif (group_fin = 1 ) then wr_ptr[odd_even_sel]++ // create the address from the write pointer and word pointer. wr_adr[odd_even_sel] = {wr_ptr[odd_even_sel],word_ptr} // concatenation 31.7.6.6 Word Count

The word count logic maintains 2 counters to track the number of words transferred from DRAM per line, one counter for odd data, and one counter for even. On receipt of a llu_go_pulse, the counters are initialized to a join_word value (number of words to the join point for that printhead channel) and the pointer select values to zero (ptr_sel). When a group of words are transferred to DRAM as indicated by the word_dec signal from the interface controller, the corresponding counter is decremented. The counter to decrement is indicated by the odd_even_sel signal from the write pointer block (even=0, odd=1).

When a counter is zero and the ptr_sel is zero, the counter is re-initialized to the second join_word value and ptr_sel is inverted. The counter continues to count down to zero each time a word_dec signal is received. When a counter is zero and the ptr_sel is one, it signals the end of a line (the last_wd signal) and initializes the counter to the first join_point value for the next line transfer.

The ptr_sel signal is used in the address generator to select the correct address pointer to use for that particular access.

// determine which counter to decrement if (llu_go_pulse == 1) then word_cnt[0] = join_word[0] // even count ptr_sel[0]  = 0 // even generator starts with pointer A word_cnt[1] = join_word[1] // odd count ptr_sel[1] = 0 // odd generator starts with pointer A elsif (word_dec == 1) then { // need to decrement one word counter if (odd_even_sel == 0) then  // even counter update if (word_cnt[0] == 0) then word_cnt[0] = join_word[ptr_sel[0]] // re-initialize pointer ptr_sel[0] = ~(ptr_sel[0]) if (ptr_sel[0]== 1) then  // determine if this the last word last_wd = 1 else word_cnt [0] −−  // normal decrement else  // odd counter update if (word_cnt[1] == 0) then word_cnt[1]  = join_word[ptr_sel[1]] // re-initialize pointer ptr_sel[1] = ~(ptr_sel[1]) if (ptr_sel[1] == 1) then  // determine if this the last word last_wd = 1 else word_cnt [1] −−  // normal decrement }

The word count logic also determines if the current word to be transferred is the join word, and if so it determines if it is aligned on a 256-bit boundary or not. If the join point is aligned to a boundary there is no need to prevent the address counter from incrementing, otherwise the address pointers are stalled for that word transfer (join_stall).

join_stall = (((ptr_sel[0] == 0)AND (word_cnt[0] == 0)AND (join_point[0][7:0] != 0)) AND ((ptr_sel[1] == 0)AND (word_cnt[1] == 0)AND (join_point[1][7:0] != 0)))

The word count logic also determines when a complete line has been read from DRAM, it then signals the fifo fill level logic in both the LLU and DWU (via line_rd signal) that a complete line has been read by the LLU (llu_dwu_line_rd).

// line finish logic if (llu_go_pulse == 1) then line_fin = 0 line_rd = 0 elsif ((last_wd == 1) AND (line_fin == 0)) then line_fin = 1  // first group last_wd finish pulse line_rd = 0 elsif ((last_wd == 1) AND (line_fin == 1)) then line_fin = 0 // second group last_wd finish pulse line_rd = 1 else line_fin = line_fin // stay the same line_rd = 0 32 Printhead Interface (PHI) 32.1 Overview

The Printhead interface (PHI) accepts dot data from the LLU and transmits the dot data to the printhead, using the printhead interface mechanism. The PHI generates the control and timing signals necessary to load and drive the bi-lithic printhead. The CPU determines the line update rate to the printhead and adjusts the line sync frequency to produce the maximum print speed to account for the printhead IC's size ratio and inherent latencies in the syncing system across multiple SoPECs.

The PHI also needs to consider the order in which dot data is loaded in the printhead. This is dependent on the construction of the printhead and the relative sizes of printhead ICs used to create the printhead. See Bi-lithic Printhead Reference document for a complete description of printhead types [10].

The printing process is a real-time process. Once the printing process has started, the next printline's data must be transferred to the printhead before the next line sync pulse is received by the printhead. Otherwise the printing process will terminate with a buffer underrun error.

The PHI can be configured to drive a single printhead IC with or without synchronization to other SoPECs. For example the PHI could drive a single IC printhead (i.e. a printhead constucted with one IC only), or dual IC printhead with one SoPEC device driving each printhead IC.

The PHI interface provides a mechanism for the CPU to directly control the PHI interface pins, allowing the CPU to access the bi-lithic printhead to:

-   -   determine printhead temperature     -   test for and determine dead nozzles for each printhead IC     -   initialize each printhead IC     -   pre-heat each printhead IC

FIG. 277 shows a high level data flow diagram of the PHI in context.

32.2 Printhead Modes of Operation

The printhead has 8 different modes of operations (although some modes are re-used). The mode of operation is defined by the state of the output pins phi_lsyncl and phi_readl and the internal printhead mode register. The modes of operation are defined in Table 210.

TABLE 210 Printhead modes of operation Internal Name Mode phi_readl phi_lsyncl State Description NORMAL XXX 1 1 N/A Normal print mode, dot data is clocked into the printhead shift register, on each falling edge of phi_srclk DOT_LOAD/ XXX 1 0 phi_frclk=0 Dot Load Mode, data stored in the FIRE_INIT dot shift register is transferred into the dot latch on the falling edge of phi_lsyncl, and latched in on the rising edge of phi_lsyncl phi_srclk=1 Fire load mode. Parameter for generating fire pattern are loaded into generator, data on phi_ph_data[1:0][0] is clocked into the generator on each rising edge of phi_frclk NOZZLE_RESET 001 0 1 N/A Reset Nozzle Test mode. Reset the state on nozzle test. CMOS_TEST 111 0 1 N/A CMOS test mode. FIRE_GEN 000 0 1 N/A Fire Initialise mode. The initialised generator creates the fire pattern and shift select pattern. The pattern is clocked into the fire shift register and select shift register on the rising edge of phi_frclk TEMP_TEST 010 0 0 N/A Temperature test output. NOZZLE_TEST 001 0 0 N/A Nozzle test output. The result of a nozzle test is output on phi_frclk_i. 32.3 Data Rate Equalization

The LLU can generate dot data at the rate of 12 bits per cycle, where a cycle is at the system clock frequency. In order to achieve the target print rate of 30 sheets per minute, the printhead needs to print a line every 100 μs (calculated from 300 mm @ 65.2 dots/mm divided by 2 seconds=˜100 μsec). For a 7:3 constructed printhead this means that 9744 cycles at 320 Mhz is quick enough to transfer the 6-bit dot data (at 2 bits per cycle). The input FIFOs are used to de-couple the read and write clock domains as well as provide for differences between consume and fill rates of the PHI and LLU.

Nominally the system clock (pclk) is run at 160 Mhz and the printhead interface clock (doclk) is at 320 Mhz.

If the PHI was to transfer data at the full printhead interface rate, the transfer of data to the shorter printhead IC would be completed sooner than the longer printhead IC. While in itself this isn't an issue it requires that the LLU be able to supply data at the maximum rate for short duration, this requires uneven bursty access to DRAM which is undesirable. To smooth the LLU DRAM access requirements over time the PHI transfers dot data to the printhead at a pre-programmed rate, proportional to the ratio of the shorter to longer printhead ICs.

The printhead data rate equalization is controlled by PrintheadRate[1:0] registers (one per printhead IC). The register is a 16 bit bitmap of active clock cycles in a 16 clock cycle window. For example if the register is set to 0xFFFF then the output rate to the printhead will be full rate, if it's set to 0xF0F0 then the output rate is 50% where there is 4 active cycles followed by 4 inactive cycles and so on. If the register was set to 0x0000 the rate would be 0%. The relative data transfer rate of the printhead can be varied from 0-100% with a granularity of 1/16 steps.

TABLE 211 Example rate equalization values for common printheads Printhead Printhead A Printhead B Ratio A:B rate (%) rate (%) 8:2 0xFFFF (100%) 0x1111 (25%) 7:3 0xFFFF (100%) 0x5551 (43.7%) 6:4 0xFFFF (100%) 0xF1F2 (68.7%) 5:5 0xFFFF (100%) 0xFFFF (100%)

If both printhead ICs are the same size (e.g. a 5:5 printhead) it may be desirable to reduce the data rate to both printhead ICs, to reduce the read bandwidth from the DRAM.

32.4 Dot Generate and Transmit Order

Several printhead types and arrangements exists (see [10] for other arrangements). The PHI is capable of driving all possible configurations, but for the purposes of simplicity only one arrangement (arrangement 1- see [10] for definition) is described in the following examples. The structure of the printhead ICs dictate the dot transmit order to each printhead IC. The PHI accepts two streams of dot data from the LLU, one even stream the other odd. The PHI constructs the dot transmit order streams from the dot generate order received from the LLU. Each stream of data has already been arranged in increasing or decreasing dot order sense by the DWU. The exact sense choice is dependent on the type of printhead ICs used to construct the printhead, but regardless of configuration the odd and even stream should be of opposing sense. The dot transmit order is shown in FIG. 281. Dot data is shifted into the printhead in the direction of the arrow, so from the diagram (taking the type 0 printhead IC) even dot data is transferred in increasing order to the mid point first (0, 2, 4, . . . , m−6, m−4, m−2), then odd dot data in decreasing order is transferred (m−1, m−3, m−5, . . . , 5, 3, 1). For the type 1 printhead IC the order is reversed, with odd dots in increasing order transmitted first, followed by even dot data in decreasing order. Note for any given color the odd and even dot data transferred to the printhead ICs are from different dot lines, in the example in the diagram they are separated by 5 dot lines. Table 212 shows the transmit dot order for some common A4 printheads. Different type printheads may have the sense reversed and may have an odd before even transmit order or vice versa.

TABLE 212 Example printhead ICs, and dot data transmit order for A4 (13824 dots) page Size Dots Dot Order Type 0 Printhead IC 8 11160 0, 2, 4, 8 . . . , 5579, 5577, 5575 . . . 5574, 5576, 5578 7, 5, 3, 1 7 9744 0, 2, 4, 8 . . . , 4871, 4869, 4867 . . . 4866, 4868, 4870 7, 5, 3, 1 6 8328 0, 2, 4, 8 . . . , 4163, 4161, 4159 . . . 4158, 4160, 4162 7, 5, 3, 1 5 6912 0, 2, 4, 8 . . . , 3455, 3453, 3451 . . . 3450, 3452, 3454 7, 5, 3, 1 4 5496 0, 2, 4, 8 . . . , 2847, 2845, 2843 . . . 2742, 2744, 2746 7, 5, 3, 1 3 4080 0, 2, 4, 8 . . . , 2039, 2037, 2035 . . . 2034, 2036, 2038 7, 5, 3, 1 2 2664 0, 2, 4, 8 . . . , 1331, 1329, 1327 . . . 1326, 1328, 1330 7, 5, 3, 1 Type 1 Printhead IC 8 11160 13823, 13821, 13819 . . . , 1332, 1334, 1336 . . . 1337, 1335, 1333 13818, 13820, 13822 7 9744 13823, 13821, 13819 . . . , 2040, 2042, 2044 . . . 2045, 2043, 2041 13818, 13820, 13822 6 8328 13823, 13821, 13819 . . . , 2848, 2850, 2852 . . . 2853, 2851, 2849 13818, 13820, 13822 5 6912 13823, 13821, 13819 . . . , 3456, 3458, 3460 . . . 3461, 3459, 3457 13818, 13820, 13822 4 5496 13823, 13821, 13819 . . . , 4164, 4166, 4168 . . . 4169, 4167, 4165 13818, 13820, 13822 3 4080 13823, 13821, 13819 . . . , 4872, 4874, 4876 . . . 4877, 4875, 4873 13818, 13820, 13822 2 2664 13823, 13821, 13819 . . . , 5580, 5582, 5584 . . . 5585, 5583, 5581 13818, 13820, 13822 32.4.1 Dual Printhead IC

The LLU contains 2 dot generator units. Each dot generator reads dot data from DRAM and generates a stream of dots in increasing or decreasing order. A dot generator can be configured to produce odd or even dot data streams, and the dot sense is also configurable. In FIG. 281 the odd dot generator is configured to produce odd dot data in decreasing order and the even dot generator produces dot data in increasing order. The LLU takes care of any vertical misalignment between the 2 printhead ICs, presenting the PHI with the appropriate data ready to be transmitted to the printhead.

In order to reconstruct the dot data streams from the generate order to the transmit order, the connection between the generators and transmitters needs to be switched at the mid point. At line start the odd dot generator feeds the type 1 printhead, and the even dot generator feeds the type 0 printhead. This continues until both printheads have received half the number of dots they require (defined as the mid point). The mid point is calculated from the configured printhead size registers (PrintheadSize). Once both printheads have reached the mid point, the PHI switches the connections between the dot generators and the printhead, so now the odd dot generator feeds the type 0 printhead and the even dot generator feeds the type 1 printhead. This continues until the end of the line.

It is possible that both printheads will not be the same size and as a result one dot generator may reach the mid point before the other. In such cases the quicker dot generator is stalled until both dot generators reach the mid point, the connections are switched and both dot generators are restarted.

Note that in the example shown in FIG. 281 the dot generators could generate an A4 line of data in 6912 cycles, but because of the mismatch in the printhead IC sizes the transmit time takes 9744 cycles.

32.4.2 Single Printhead IC

In some cases only one printhead IC may be connected to the PHI. In FIG. 282 the dot generate and transmit order is shown for a single IC printhead of 9744 dots width. While the example shows the printhead IC connected to channel A, either channel could be used. The LLU generates odd and even dot streams as normal, it has no knowledge of the physical printhead configuration. The PHI is configured with the printhead size (PrintheadSize[1] register) for channel B set to zero and channel A is set to 9744.

Note that in the example shown in FIG. 283 the dot generators could generate an 7 inch line of data in 4872 cycles, but because the printhead is using one IC, the transmit time takes 9744 cycles, the same speed as an A4 line with a 7:3 printhead.

32.4.3 Summary of Generate and Transmit Order Requirements

In order to support all the possible printhead arrangements, the PHI (in conjuction with the LLU/DWU) must be capable of re-ordering the bits according to the following criteria:

-   -   Be able to output the even or odd plane first.     -   Be able to output even and odd planes independently.     -   Be able to reverse the sequence in which the color planes of a         single dot are output to the printhead.         32.5 Print Sequence

The PHI is responsible for accepting dot data streams from the LLU, restructuring the dot data sequence and transferring the dot data to each printhead within a line time (i.e before the next line sync).

Before a page can be printed the printhead ICs must be initialized. The exact initialization sequence is configuration dependent, but will involve the fire pattern generation initialization and other optional steps. The initialization sequence is implemented in software.

Once the first line of data has been transferred to the printhead, the PHI will interrupt the CPU by asserting the phi_icu_print_rdy signal. The interrupt can be optionally masked in the ICU and the CPU can poll the signal via the PCU or the ICU. The CPU must wait for a print ready signal in all printing SoPECs before starting printing.

Once the CPU in the PrintMaster SoPEC is satisfied that printing should start, it triggers the LineSyncMaster SoPEC by writing to the PrintStart register of all printing SoPECs. The transition of the PrintStart register in the LineSyncMaster SoPEC will trigger the start of lsyncl pulse generation. The PrintMaster and LineSyncMaster SoPEC are not necessarily the same device, but often are the same. For a more in depth definition see section 12.1.1 Multi-SoPEC systems on page 105.

Writing a 1 to the PrintStart register enables the generation of the line sync in the LineSyncMaster which is in turn used to align all SoPECs in a multi-SoPEC system. All printhead signaling is aligned to the line sync. The PrintStart is only used to align the first line sync in a page.

When a SoPEC receives a line sync pulse it means that the line previously transferred to the printhead is now printing, so the PHI can begin to transfer the next line of data to the printhead. When the transfer is complete the PHI will wait for the next line sync pulse before repeating the cycle. If a line sync arrives before a complete line is transferred to the printhead (i.e. a buffer error) the PHI generates a buffer underrun interrupt, and halts the block.

For each line in a page the PHI must transfer a full line of data to the printhead before the next line sync is generated or received.

32.5.1 Sync Pulse Control

If the PHI is configured as the LineSyncMaster SoPEC it will start generating line sync signals LsyncPre number of pclk cycles after PrintStart register rising transition is detected. All other signals in the PHI interface are referenced from the rising edge of phi_lsyncl signal.

If the SoPEC is in line sync slave mode it will receive a line sync pulse from the LineSyncMaster SoPEC through the phi_lsyncl pin which will be programmed into input mode. The phi_lsyncl input pin is treated as an asynchronous input and is passed through a de-glitch circuit of programmable de-glitch duration (LsyncDeglitchCnt).

The phi_lsyncl will remain low for LsyncLow cycles, and then high for LsyncHigh cycles. The phi_lsyncl profile is repeated until the page is complete. The period of the phi_lsyncl is given by LsyncLow+LsyncHigh cycles. Note that the LsyncPre value is only used to vary the time between the generation of the first phi_lsyncl and the PageStart indication from the CPU. See FIG. 284 for reference diagram.

If the SoPEC device is in line sync slave mode, the LsyncHigh register specifies the minimum allowed phi_lsyncl period. Any phi_lsyncl pulses received before the LsyncHigh has expired will trigger a buffer underrun error.

32.5.2 Shift Register Signal Control

Once the PHI receives the line sync pulse, the sequence of data transfer to the printhead begins. All PHI control signals are specified from the rising edge of the line sync.

The phi_srclk (and consequently phi_ph_data) is controlled by the SrclkPre, SrclkPost registers. The SrclkPre specifies the number of pclk cycles to wait before beginning to transfer data to the printhead. Once data transfer has started, the profile of the phi_srclk is controlled by PrintheadRate register and the status of the PHI input FIFO. For example it is possible that the input FIFO could empty and no data would be transferred to the printhead while the PHI was waiting. After all the data for a printhead is transferred to the PHI, it counts SrclkPost number of pclk cycles. If a new phi_lsyncl falling edge arrives before the count is complete the PHI will generate a buffer underrun interrupt (phi_icu_underrun).

32.5.3 Firing Sequence Signal Control

The profile of the phi_frclk pulses per line is determined by 4 registers FrclkPre, FrclkLow, FrclkHigh, FrclkNum. The FrclkPre register specifies the number of cycles between line sync rising edge and the phi_frclk pulse high. It remains high for FrclkHigh cycles and then low for FrclkLow cycles. The number of pulses generated per line is determined by FrclkNum register. The total number of cycles required to complete a firing sequence should be less than the phi_lsyncl period i.e. ((FrclkHigh+FrclkLow)*FrclkNum)+FrclkPre<(LsyncLow+LsyncHigh). Note that when in CPU direct control mode (PrintheadCpuCtrl=1) and PrintheadCpuCtrlMode[x]=1, the frclk generator is triggered by the transition of the FireGenSoftTrigger[0]0 bit from 0 to 1. FIG. 284 details the timing parameters controlling the PHI. All timing parameters are measured in number of pclk cycles.

32.5.4 Page Complete

The PHI counts the number of lines processed through the interface. The line count is initialised to the PageLenLine and decrements each time a line is processed. When the line count is zero it pulses the phi icu_page_finish signal. A pulse on the phi_icu_page_finish automatically resets the PHI Go register, and can optionally cause an interrupt to the CPU. Should the page terminate abnormally, i.e. a buffer underrun, the Go register will be reset and an interrupt generated.

32.5.5 Line Sync Interrupt

The PHI will generate an-interrupt to the CPU after a predefined number of line syncs have occured. The number of line syncs to count is configured by the LineSyncInterrupt register. The interrupt can be disabled by setting the register to zero.

32.6 Dot Line Margin

The PHI block allows the generation of margins either side of the received page from the LLU block. This allows the page width used within PEP blocks to differ from the physical printhead size.

This allows SoPEC to store data for a page minus the margins, resulting in less storage requirements in the shared DRAM and reduced memory bandwidth requirements. The difference between the dot data line size and the line length generated by the PHI is the dot line margin length. There are two margins specified for any sheet, a margin per printhead IC side.

The margin value is set by programming the DotMargin register per printhead IC. It should be noted that the DotMargin register represents half the width of the actual margin (either left or right margin depending on paper flow direction). For example, if the margin in dots is 1 inch (1600 dots), then DotMargin should be set to 800. The reason for this is that the PHI only supports margin creation cases 1 and 3 described below.

See example in FIG. 284.

In the example the margin for the type 0 printhead IC is set at 100 dots (DotMargin==100), implying an actual margin of 200 dots.

If case one is used the PHI takes a total of 9744 phi_srclk cycles to load the dot data into the type 0 printhead. It also requires 9744 dots of data from the LLU which in turn gets read from the DRAM. In this case the first 100 and last 100 dots would be zero but are processed though the SoPEC system consuming memory and DRAM bandwidth at each step.

In case 2 the LLU no longer generates the margin dots, the PHI generates the zeroed out dots for the margining. The phi_srclk still needs to toggle 9744 times per line, although the LLU only needs to generate 9544 dots giving the reduction in DRAM storage and associated bandwidth.

The case 2 senario is not supported by the PHI because the same effect can be supported by means of case 1 and case 3.

If case 3 is used the benefits of case 2 are achieved, but the phi_srclk no longer needs to toggle the full 9744 clock cycles. The phi_srclk cycles count can be reduced by the margin amount (in this case 9744-100=9644 dots), and due to the reduction in phi_srclk cycles the phi_lsyncl period could also be reduced, increasing the line processing rate and consequently increasing print speed. Case 3 works by shifting the odd (or even) dots of a margin from line Y to become the even (or odd) dots of the margin for line Y-4, (Y-5 adjusted due to being printed one line later). This works for all lines with the exception of the first line where there has been no previous line to generate the zeroed out margin. This situation is handled by adding the line reset sequence to the printhead initialization procedure, and is repeated between pages of a document.

32.7 Dot Counter

For each color the PHI keeps a dot usage count for each of the color planes (called AccumDotCount). If a dot is used in particular color plane the corresponding counter is incremented. Each counter is 32 bits wide and saturates if not reset. A write to the DotCountSnap register causes the AccumDotCount[N] values to be transferred to the DotCount[N] registers (where N is 5 to 0, one per color). The AccumDotCount registers are cleared on value transfer.

The DotCount[N] registers can be written to or read from by the CPU at any time. On reset the counters are reset to zero.

The dot counter only counts dots that are passed from the LLU through the PHI to the printhead.

Any dots generated by direct CPU control of the PHI pins will not be counted.

32.8 CPU IO Control

The PHI interface provides a mechanism for the CPU to directly control the PHI interface pins, allowing the CPU to access the bi-lithic printhead:

-   -   Determine printhead temperature     -   Test for and determine dead nozzles for each printhead IC     -   Printhead IC initialization     -   Printhead pre-heat function

The CPU can gain direct control of the printhead interface connections by setting the PrintheadCpuCtrl register to one. Once enabled the printhead bits are driven directly by the PrintheadCpuOut control register, where the values in the register are reflected directly on the printhead pins and the status of the printhead input pins can be read directly from the PrintheadCpuIn. The direction of pins is controlled by programming PrintheadCpuDir register.

The register to pin mapping is as follows:

TABLE 213 CPU control and status registers mapping to printhead interface Register Name bits Printhead pin PrintHeadCpuOut 0 phi_lsyncl_o 1 phi_frclk_o 2 Reserved 4:3 phi_ph_data_o[0][1:0] 6:5 phi_ph_data_o[1][1:0] 8:7 phi_srclk[1:0] 9 phi_readl PrintHeadCpuDir 0 phi_lsyncl_e direction control 1 - output mode 0 - input mode 1 phi_frclk_e direction control 1 - output mode 0 - input mode 2 Reserved PrintHeadCpuIn 0 phi_lsyncl_i 1 phi_frclk_i 2 Reserved

It is important to note that once in PrintheadCpuCtrl mode it is the responsibility of the CPU to drive the printhead correctly and not create situations where the printhead could be destroyed such as activating all nozzles together.

The phi_srclk is a double data rate clock (DDR) and as such will clock data on both edges in the printhead.

Note the following procedures are based on current printhead capabilities, and are subject to change.

32.9 Implementation

32.9.1 Definitions of I/O

TABLE 214 Printhead interface I/O definition Port name Pins I/O Description Clocks and Resets Pclk 1 In System Clock Doclk 1 In Data out clock (2 × pclk) used to transfer data to printhead prst_n 1 In System reset, synchronous active low. Synchronous to pclk dorst_n 1 In System reset, synchronous active low. Synchronous to doclk General phi_icu_print_rdy 1 Out Indicates that the first line of data is transferred to the printhead Active high. phi_icu_page_finish 1 Out Indicates that data for a complete page has transferred. Active high phi_icu_underrun 1 Out Indicates the PHI has detected a buffer underrun. Active high phi_icu_linesync_int 1 Out Indicates the PHI has detected LineSyncInterrupt number of line syncs. Debug debug_data_valid 1 In Output debug data valid to be muxed on to the PHI pin debug_cntrl 1 In Control signal for the PHI to indicate whether or not the debug data valid (and pclk) should be selected by the pin mux. Active high. LLU Interface llu_phi_data[1:0][5:0] 2 × 6 In Dot Data from LLU to the PHI, each bit is a color plane 5 downto 0. Bus 0 - Even dot data stream Bus 1 - Odd dot data stream Data is active when corresponding bit is active in llu_phi_avail bus phi_llu_ready[1:0] 2 Out Indicates that PHI is ready to accept data from the LLU 0 - Even dot data stream 1 - Odd dot data stream llu_phi_avail[1:0] 2 In Indicates valid data present on corresponding llu_phi_data. 0 - Even dot data stream 1 - Odd dot data stream Printhead Interface phi_ph_data[1:0][1:0] 2 × 2 Out Dot data output to printhead. Each bus to each printhead contains 2 bits of data Bus 0 - Printhead channel A Bus 1 - Printhead channel B phi_srclk[1:0] 2 Out Dot data shift clock used to clock in printhead data, data is shifted on both edges of clock(i.e. double data rate DDR). Bus 0 - Printhead channel A Bus 1 - Printhead channel B phi_readl 1 Out Common printhead mode control. Used in conjunction with phi_lsyncl to determine the printhead mode 0 - SoPEC receiving, printhead driving 1 - SoPEC driving, printhead receiving phi_frclk_o 1 Out Common Fire pattern clock needs to toggle once per fire cycle phi_frclk_e 1 In phi_frclk_o output enable, when high phi_frclk_o pin is driving phi_frclk_l 1 In phi_frclk_i input from printhead phi_lsyncl_o 1 Out Capture dot data for next print line, output mode phi_lsyncl_e 1 In phi_lsyncl output enable, when high phi_lsyncl pin is driving phi_lsyncl_i 1 In Line Sync Pulse from Master SoPEC PCU Interface pcu_phi_sel 1 In Block select from the PCU. When pcu_phi_sel is high both pcu_adr and pcu_dataout are valid. pcu_rwn 1 In Common read/not-write signal from the PCU. pcu_adr[7:2] 6 In PCU address bus. Only 6 bits are required to decode the address space for this block. pcu_dataout[31:0] 32 In Shared write data bus from the PCU. phi_pcu_rdy 1 Out Ready signal to the PCU. When phi_pcu_rdy is high it indicates the last cycle of the access. For a write cycle this means pcu_dataout has been registered by the block and for a read cycle this means the data on phi_pcu_datain is valid. phi_pcu_datain[31:0] 32 Out Read data bus to the PCU. 32.9.2 PHI Sub-Block Partition 32.9.3 Configuration Registers

The configuration registers in the PHI are programmed via the PCU interface. Refer to section 21.8.2 on page 321 for a description of the protocol and timing diagrams for reading and writing registers in the PHI. Note that since addresses in SoPEC are byte aligned and the PCU only supports 32-bit register reads and writes, the lower 2 bits of the PCU address bus are not required to decode the address space for the PHI. When reading a register that is less than 32 bits wide zeros should be returned on the upper unused bit(s) of phi_pcu_datain. Table 215 lists the configuration registers in the PHI

TABLE 215 PHI registers description Address PHI_base+ Register #bits Reset Description Control Registers 0x00 Reset 1 0x1 Active low synchronous reset, self de- activating. A write to this register will cause a PHI block reset. 0x04 Go 1 0x0 Active high bit indicating the PHI is programmed and ready to use. A low to high transition will cause PHI block internal state to reset. Will be automatically reset if a page finish or a buffer underrun is detected. General Control 0x08 PageLenLine 32 0x0000_0000 Specifies the number of dot lines in a page. Indicates the number of lines left to process in this page while the PHI is running (Working register) 0x0c PrintStart 1 0x0 A high level enables printing to start via the generation of line syncs in a master, and acceptance of line syncs in a slave. Can be set in advance of the print ready signal. 0x10-0x14 DotMargin[1:0] 2 × 16 0x0000 Specifies for each printhead IC, the width of the margin in dots divided by 2. Value must be divisible by 2 (i.e. the low bit must be 0) 0 - Printhead IC Channel A 1 - Printhead IC Channel B 0x18-0x2C DotCount[5:0] 6 × 32 0x0000_0000 Indicates the number of Dots used for a particular color, where N specifies a color from 0 to 5. Value valid after a write access to DotCountSnap 0x30 DotCountSnap 1 0x0 Write access causes the AccumDotCount values to be transferred to the DotCount registers. The AccumDotCount are reset afterwards.(Reads as zero) 0x34 PhiHeadSwap 1 0x0 Controls which signals are connected to printhead channels A and B 0 - Normal, specifies bit 0 is channel A, bit 1 is channel B 1 - Swapped, specifies bit 0 is channel B, bit 1 is channel A. 0x38 PhiMode 1 0x0 Indicates whether the PHI is operating in master or slave mode 0 - Slave Mode 1 - Master Mode 0x3C-0x40 PhiSerialOrder 2 × 1  0x0 Specifies the serialization order of dots before transfer to the printhead. Bus 0 - Printhead Channel A Bus 1 - Printhead Channel B If set to zero the order is dot[1:0], then dot[3:2] then dot[5:4]. If set to one then the order is dot[5:4], dot[3:2], dot[1:0]. 0x44-0x48 PrintHeadSize 2 × 16 0x0000 Specifies the number of non-margin dots in the printhead ICs (must be even). If margining is to be used then the configured PrintHeadSize should be adjusted by the dot margin value i.e. PrintHeadSize = (Physical- PrintHeadSize − (DotMargin * 2)). Value must be divisible by 2 (i.e. the low bit must be 0) Bus 0 - Specifies printhead on Channel A Bus 1 - Specifies printhead on Channel B CPU Direct PHI Control (See Table 213.) 0x4C PrintHeadCpuln 3 0x0 PHI interface pins input status. Only active in direct CPU mode (Read Only Register) 0x50 PrintHeadCpuDir 3 0x0 PHI interface pins direction control. Only active in direct CPU mode 0x54 PrintHeadCpuOut 10 0x000 PHI interface pins output control. Only active in direct CPU mode 0x58 PrintHeadCpuCtrl 1 0x1 Control direct access CPU access to the PHI pins 0 - Normal Mode 1 - Direct CPU Control mode 0x5C Print-HeadCpuCtrlMode 1 0x0 Specifies if the pin is controlled by the PrintHeadCpuOut register or by the Fire generator logic. Only active when PrintHeadCpuCtrl is 1 and pin is in output mode. Bit 0 - controls the frclk pin When the bit is 0 - Pin is controlled by PrintHeadCpuOut 1 - Pin is controlled by Fire Generator Logic Line Sync Control 0x60 LsyncHigh 24 0x00_0000 In Master mode specifies the number of pclk cycles phi_lsyncl should remain high. In Slave mode specifies the minimum number of pclk cycles between Lsync pulses. Lsync pulses of a shorter period will cause the PHI to halt due to buffer underrun. 0x64 LsyncLow 16 0x0000 Number of pclk cycles phi_lsyncl should remain low. 0x68 LsyncPre 16 0x0000 Number of pclk cycles between PrintStart rising transition and the generated phi_lsyncl falling edge 0x6C LsyncDeglitchCnt 4 0x3 Number of pclk cycles to filter the incoming Lsync pulse from the master. Only used in slave mode. 0x70 LineSyncInterrupt 16 0x0000 Number of line syncs to occur before generating an interrupt. When set to zero interrupt is disabled. Shift Register Control 0x74 SrclkPre 14 0x0000 Number of pclk cycles between phi_lsyncl falling edge and phi_srclk pulse generation, or printhead data transfer 0x78 SrclkPost 14 0x0000 Number of pclk cycles allowed margin from last srclk pulse in a line to before next line sync 0x7C-0x80 PrintHeadRate[1:0] 2 × 16 0xFFFF Specifies the active to inactive ratio of phi_srclk for the printhead ICs. A 1 indicates Active. Bus 0 - Printhead IC channel A Bus 1 - Printhead IC channel B 0x84 DotOrderMode 1 0x0 Specifies the dot transmit order to the printhead Channel A. Printhead Channel B is always the opposing order. 0 - Even before Odd dots 1 - Odd before Even dots Fire Control 0x98 FrclkPre 14 0x0000 Number of pclk cycles after lsyncl transitions from 0 to 1 to phi_frclk pulse generation 0x9C FrclkLow 14 0x0000 Number of pclk cycles phi_frclk should remain low. 0xA0 FrclkHigh 14 0x0000 Number of pclk cycles phi_frclk should remain high. 0xA4 FrclkNum 16 0x0000 Number of phi_frclk pulses per line time. 0xA8 FireGenSoftTrigger 1 0x0 Only active when PrintHeadCpuCtrlMode is set to 1, PrintHeadCpuCtrl is 1 and pin is in output mode. Bit 0 controls frclk generator. A 0 to 1 transition on a bit triggers the corresponding generator to create the programmed pulse profile (configured by FrclkNum, FrclkHigh, FrclkLow, FrclkPre registers) when complete the bit gets reset to 0. Working Registers 0xAC-0xB0 LineDotCnt 2 × 16 0x0000 Indicates the number of dot processed in the current line Bus 0 - Printhead Channel A Bus 1 - Printhead Channel B (Read Only Registers)

The configuration registers in the PHI block are clocked at pclk rates but some blocks in the PHI are clocked by different and asynchronous clocks. Configuration values are not re-synchronized, it is therefore important that the Go register be set to zero while updating configuration values. This prevents logic from entering unknown states due to metastable clock domain transfers.

Some registers can be written to at any time such as the direct CPU control registers (PrintheadCpuIn, PrintheadCpuDir, PrintheadCpuOut and PrintheadCpuCtrl), the Go register and the PrintStart register. All registers can be read from at any time.

32.9.4 Dot Counter

The dot counter keeps a running count of the number of dots fired for each color plane. The counters are 32 bits wide and will saturate. When the CPU wants to read the dot count for a particular color plane it must write to the DotCountSnap register. This causes all 6 running counter values to be transferred to the DotCount registers in the configuration registers block. The running counter values are reset.

// reset if being snapped if (dot_cnt_snap == 1) then{ dot_count[5:0] = accum_dot_count[5:0] accum_dot_count[5:0] = 0 } // update the counts for (color=0;color < 6;color++) { if (accum_dot_count[color] != 0xffff_ffff) { // data valid, first dot stream data_valid  = ((phi_llu_ready[0] == 1) AND (llu_phi_avail[0] == 1)) if ((data_valid == 1) AND (llu_phi_data[0][color] == 1)) then accum_dot_count[color] ++ // data valid, second dot stream data_valid  =  ((phi_llu_ready[1]  ==  1)  AND (llu_phi_avail[1] == 1)) if ((data_valid == 1) AND (llu_phi_data[1][color] == 1)) then accum_dot_count[color] ++ } } 32.9.5 Sync Generator

The sync generator logic has two modes of operation, master and slave mode. In master mode (configured by the PhiMode register) it generates the lsyncl_o output based on configured values and control triggers from the PHI controller. In slave mode it de-glitches the incoming lsyncl_i signal, and filters the lsyncl signal with the minimum configured period.

After reset or a pulse on phi_go_pulse the machine returns to the Reset state, regardless of what state it's currently in.

The state machine waits until it's enabled (sync_en==1) by the PHI controller state machine. When enabled it can proceed to the SyncPre or SyncWait depending on whether the state machine is configured in master or slave mode. In master mode it generates the lsyncl pulses, in slave mode it receives and filters the lsyncl pulses from the master sync generator.

On transition to the SyncPre state a counter is loaded with the LsyncPre value, and while in the SyncPre the counter is decremented. When the count is zero the machine proceeds to the SyncLow state loading the counter with LsyncLow value.

The machine waits in the SyncLow state until the counter has decremented to zero. It proceeds to the SyncHigh state pulsing the line_st signal on transition and counts LsyncHigh number of cycles. This indicates to the PHI controller the line start aligned to the lsyncl positive edge. While in LsyncLow state the lsyncl_o output is set to 0 and in SyncHigh the lsyncl_o output is set to 1. When the count is zero and the current line is not the last (last_line==0), the machine returns to the SyncLow state to begin generating a new line sync pulse. The transition pulses the line_fin signal to the PHI controller.

The loop is repeated until the current line is the last (last_line==1), and the machine returns to the Reset state to wait for the next page start.

In slave mode the state machine proceeds to the SyncWait state when enabled. It waits in this state until a-lsync_pulse_rise is received from the input de-glitch circuit. When a pulse is detected the machine jumps to the SyncPeriod state and begins counting down the LsyncHigh number of clock cycles before returning to the SyncWait state. Note in slave mode the LsyncHigh specifies the minimum number of pclk cycles between Lsync pulses. On transition from the SyncWait to the SyncPeriod state the line_st signal to the PHI controller is pulsed to indicate the line start. While in the SyncPeriod state if a lsync_pulse_fall is detected the state machine will signal a sync error (via sync_err) to the PHI controller and cause a buffer underrun interrupt.

32.9.5.1 Lsyncl Input De-Glitch

The lsync_i input is considered an asynchronous input to the PHI, and is passed through a synchronizer to reduce the possibility of metastable states occurring before being passed to the de-glitch logic.

The input de-glitch logic rejects input states of duration less than the configured number of clock cycles (lsync_deglitch_cnt), input states of greater duration are reflected on the output, and are negative and positive edge detected to produce the lsync_pulse_fall and lsync_pulse_rise signal to the main generator state machine. The counter logic is given by

if ( lsync_i != lsync_i_delay) then cnt = lsync_deglitch_cnt output_en = 0 elsif (cnt == 0 ) then cnt = cnt output_en = 1 else cnt −− output_en = 0 32.9.5.2 Line Sync Interrupt Logic

The line sync interrupt logic counts the number of line syncs that occur (either internally or externally generated line syncs) and determines whether to generate an interrupt or not. The number of line syncs it counts before an interrupt is generated is configured by the LineSyncInterrupt register. The interrupt is disabled if LineSyncInterrupt is set to zero.

// implement the interrupt counter if (phi_go_pulse ==1) then line_count = 0 elsif (line_st == 1) AND (line_count == 0)) then line_count = linecount_int elsif ((line_st == 1) AND (line_count != 0)) then line_count −− // determine when to pulse the interrupt if (linesync_int == 0 ) then // interrupt disabled phi_icu_linesync_int = 0; elsif ((line_st == 1) AND (line_count == 1)) then phi_icu_linesync_int = 1 32.9.6 Fire Generator

The fire generator block creates the signal profile for the phi_frclk signal to the printhead. The frclk is based on configured values and is timed in relation to the fire_st pulse from the PHI controller block. Should the phi_frclk state machine receive a fire_st pulse before it has completed the sequence the machine will restart regardless of its current state.

Alternatively the frclk state machine can be triggered to generate their configured pulse profile by software. A low to high transition on the FireGenSoftTrigger register will cause a pulse on soft_frclk_st triggering the state machine to begin generating the pulse profile. When the state machine has completed its sequence it will clear the FireGenSoftTrigger register bit (via soft_fire_clr signal). The FireGenSoftTrigger register will only be active when the printhead interface is in CPU direct control mode (PrintheadCpuCtrl=1), the fire generator is in software trigger mode (PrintheadCpuCtrlMode[x]=1) and the pin is configured to be output mode (PrintheadCpuDir[x]=1).

The fire generator consists of a state machine for creating the phi_frclk signal. The phi_frclk signal is generated relative to the lsyncl signal.

The machine is reset to the Reset state when phi_go_pulse==1 or the reset is active, regardless of the current state.

The machine waits in the reset state until it receives a fire_st pulse from the PHI controller (or an soft_fire_st from the configuration registers). The controller will generate a fire_st pulse at the beginning of each dot line. On the state transition the cycle counter is loaded with the FrclkPre value and the repeat counter is loaded with the FrclkNum value.

The state machine waits in the FirePre state until the cycle counter is zero, after which it jumps to the FireHigh state and loads the cycle counter with FrclkHigh value. Again the state machine waits until the count is zero and then proceeds to the FireLow state. On transition the cycle counter is loaded with the FireLow value. The state machine waits in the FireLow state while the cycle counter is decremented.

When the cycle counter reaches zero and the repeat_count is non-zero, the repeat_count is decremented, the cycle counter is loaded with the FrclkHigh value and the state machine jumps to the FireHigh state to repeat the phi_frclk generation cycle. The loop is repeated until the repeat_count is zero. In such cases_the state machine goes to the reset state resetting FireGenSoftTrigger (via the soft_fire_clr signal) register on the transition and waits for the next fire_st pulse.

When in the Reset state the fire_rdy signal is active to indicate to the controller that the fire generator is ready.

32.9.7 PHI Controller

The PHI controller is responsible for controlling all functions of the PHI block on a line by line basis. It controls and synchronizes the sync generator, the fire generator, and datapath unit, as well as signalling back to the CPU the PHI status. It also contains a line counter to determine when a full page has completed printing.

The PHI controller state machine is reset to Reset state by a reset or phi_go_pulse==1.

It will remain in reset until the block is enabled by phi_go==1. Once enabled the state machine will jump to the FirstLine state, trigger the transfer of one line of data to the printhead (data_st==1) and the line counter will be initialized to the page length (PageLenLine). Once the line is transferred (data_fin from the datapath unit) the machine will go to PrintStart state and signal the CPU using an interrupt that the PHI is ready to begin printing (phi_icu_print_rdy). The line counter will also be decremented. It will then wait in the PrintStart state until the CPU acknowledges the print ready signal and enables printing by writing to the PrintStart register.

The state machine proceeds to the SyncWait state and waits for a line start condition (line_st==1). The line start condition is different depending on whether the PHI is configured as being in a master or slave SoPEC (the PhiMode register). In either case the sync generator determines the correct line start source and signals the PHI controller via the line_st signal. Once received the machine proceeds to the LineTrans state, with the transition triggering the fire generator to start (fire_st), the datapath unit to start (data_st) and the sync generator to start (sync_st).

While in the LineTrans state the fire, sync and datapath unit will be producing line data. When finished processing a line the datapath unit will assert the line finished (data_fin) signal. If the line counter is not equal to 1 (i.e. not the last line) the state machine will jump back to the SyncWait state and wait for the start condition for the next line. The line counter will be decremented. If the line counter is one then the machine will proceed to the LastLine state.

The LastLine state generates one more line of fire pulses to print the last line held in the shift registers of the printhead. Once complete (fire_fin==1) the state machine returns to the reset state and waits for the next page of data. On page completion the state machine generates a phi_icu_page_finish interrupt to signal to the CPU that the page has completed, the phi_icu_page_finish will also cause the Go register to reset automatically.

While the state machine is in the LineTrans state (or in FirstLine state and the PHI is in slave mode) and waiting for the datapath unit to complete line processing, it is possible (e.g. an excessive PEP stall) that a line finish condition occurs (line_fin==1) but the datapath unit is not ready. In this case an underrun error is generated. The state machine goes to the Underrun state and generates a phi_icu_underrun interrupt to the CPU. The PHI cannot recover from a buffer underrun error, the CPU must reset the PEP blocks and re-start printing. The phi_icu_underrun will also cause the Go register to reset automatically.

32.9.8 CPU IO Control

The CPU IO control block is responsible for providing direct CPU control of the IO pins via the configuration registers. It also accepts the input signals from the printhead and re-synchronizes them to the pclk domain, and debug signals from the RDU and muxes them to output pins. Table contains the direct mapping of configuration registers to printhead IO pins. Direct CPU control is enabled only when PrintheadCpuCtrl is set to one. In normal operation (i.e. PrintheadCpuCtrl==0) the printhead frclk pin is always in output mode (phi_frclk_e=1), the phi_lsyncl will be in output if the SoPEC is the master, i.e. phi_lsyncl_e=phi_mode, and readl will be set high.

The PrintheadCpuCtrlMode register determine whether the frclk pin should be driven by the fire generator logic or direct from the CPU PrintheadCpuOut register.

The pseudocode for the CPU IO control is:

if (printhead_cpu_ctrl == 1) then // CPU access enabled // outputs if (PrintHeadCpuCtrlMode[0] == 1) then  // fire generator controlled phi_frclk_o = frclk else // normal direct CPU control phi_frclk_o = printhead_cpu_out[1] phi_ph_data_o[0][1:0]  = printhead_cpu_out[4:3] phi_ph_data_o[1][1:0]  = printhead_cpu_out[6:5] phi_srclk[1:0]  = printhead_cpu_out[8:7] phi_readl  = printhead_cpu_out[9] // direction control phi_lsyncl_e  = printhead_cpu_dir[0] phi_frclk_e  = printhead_cpu_dir[1] // input assignments printhead_cpu_in[0]  = synchronize(phi_lsyncl_i) printhead_cpu_in[1]  = synchronize(phi_frclk_i) else // normal connections // outputs phi_ph_data_o[0][1:0] = ph_data[0][1:0] phi_ph_data_o[1][1:0] = ph_data[1][1:0] phi_lsyncl_o = lsync_o phi_readl = 1 phi_srclk[1:0] = srclk[1:0] phi_frclk_o = frclk // direction control phi_frclk_e = 1 phi_lsyncl_e  = phi_mode // depends on Master or Slave mode // inputs lsyncl_i = phi_lsync_i // connected regardless // debug overrides any other connections if (debug_cntrl[0] == 1) then phi_frclk_o = debug_data_valid phi_frclk_e = 1 phi_readl = pclk

The debug signalling is controlled by the RDU block (see Section 11.8 Realtime Debug Unit (RDU)), the 10 control in the PHI muxes debug data onto the PHI pins based on the control signals from the RDU.

32.9.9 Datapath Unit

32.9.10 Dot Order Controller

The dot order controller is responsible for controlling the dot order blocks. It monitors the status of each block and determines the switch over point, at which the connections from odd and even dot streams to printhead channels are swapped.

The machine is reset to the Reset state when phi_go_pulse==1 or the reset is active. The machine will wait until it receives a data st pulse from the PHI controller before proceeding to the LineStart state. On the transition to the LineStart state it will reset the dot counter in each dot order block via the dot_cnt_rst signal.

While in the LineStart state both dot order blocks are enabled (gen_en==1). The dot order blocks process data until each of them reach their mid point. The mid point of a line is defined by the configured printhead size (i.e. print_head_size). When a dot order block reaches the mid point it immediately stops processing and waits for the remaining dot order block. When both dot order blocks are at the mid point (mid_pt==11) the controller clocks through the LineMid state to allow the pipeline to empty and immediately goes to LineEnd state.

In the LineEnd state the mode_sel is switched and the dot order blocks re-enabled, in this state the dot order blocks are reading data from the opposite LLU dot data stream as in LineStart state. The controller remains in the LineEnd state until both dot order blocks have processed a line i.e. line_fin==11.

On completion of both blocks the controller returns to the Reset state and again awaits the next data_st pulse from the PHI controller. When in Reset state the machine signals the PHI controller that it's ready to begin processing dot data via the dot_order_rdy signal.

The dot order controller selects which dot streams should feed which printhead channels. The order can be changed by configuring the DotOrderMode register. In all cases Channel A and Channel B must be in opposing dot order modes. Table 216 shows the possible modes of operation.

TABLE 216 Mode selection in Dot order controller. Channel Mode_sel DotOrderMode Dot transmit order A 0 0 Even before Odd (EBO mode), even dot stream feeds Channel A printhead, first half line. 0 1 Odd before Even (OBE mode), odd dot stream feeds Channel A printhead, first half line. 1 0 Even before Odd (EBO mode), even dot stream feeds Channel A printhead, second half line. 1 1 Odd before Even (OBE mode), odd dot stream feeds Channel A printhead, second half line. B 0 0 Odd before Even (OBE mode), odd dot stream feeds Channel B printhead, second half line 0 1 Even before Odd (EBO mode), even dot stream feeds Channel B printhead, second half line. 1 0 Odd before Even (OBE mode), odd dot stream feeds Channel B printhead, first half line. 1 1 Even before Odd (EBO mode), even dot stream feeds Channel B printhead, first half line. 32.9.10.1 Dot Order Unit

The dot order control accepts dot data from either dot stream from the LLU and writes the dot data into the dot buffer. It has two modes of operation, odd before even (OBE) and even before odd (EBO). In the OBE mode data from the odd stream dot data is accepted first then even, in EBO mode it's vice versa. The mode is configurable by the DotOrderMode register.

The dot order unit maintains a dot count that is decremented each time a new dot is received from the LLU. The dot order controller resets the dot counter to the print_head_size[15:0] at the start of a new line via the dot_cnt_rst signal. The dot count is compared with the printhead size (print_head_size[15:0] divided by 2) to determine the mid point (mid_Pt) and the line finish point (line_fin) when the dot counter is zero.

The mid point is defined as the half the number of dots in a particular printhead, and is derived from the the print_head_size bus by dividing by 2 and rounding down.

// define the mid point if (dot_cnt[15:0] == print_head_size[15:1] )then mid_pt = 1 else mid_pt = 0

The dot order unit logic maintains the dot data write pointer. Each time a new dot is written to the dot buffer the write pointer is incremented. The fill level of the dot buffer is determined by comparing the read and write pointers. The fill level is used to determine when to backpressure the LLU (ready signal) due to the dot buffer filling. A suitable threshold value is determined to allow for the full LLU pipeline to empty into the dot buffer.

The dot order stalling control is given by:

// determine the ready/avail signal to use, based on mode select if (mode_sel == 1) then dot_active = llu_phi_avail[0] AND ready wr_data = llu_phi_data[0] else dot_active = llu_phi_avail[1] AND ready wr_data = llu_phi_data[1] // update the counters if (dot_active == 1) then { wr_en = 1 wr_adr ++ if (dot_cnt == 0) then dot_cnt = print_head_size else dot_cnt−− }

The dot writer needs to determine when to stall the LLU dot data stream. A number of factors could stall the dot stream in the LLU such as buffer filling, waiting for the mid point, waiting for the line finish or the dot order controller is waiting for the line start condition from the PHI controller.

The stall logic is given by:

// determine when to stall the LLU generator fill_level = wr_adr − rd_adr if (fill_level > (32 − THRESHOLD ))then // THRESHOLD is open value ready = 0 // buffer is close to full elsif ( gen_en == 0) then ready = 0  // stalled by the datapath controller else ready = 1 // everything good no stall 32.9.10.2 Data Generator

The data generator block reads data from the dot buffer and feeds dot data to the printhead at a configured rate (set by the PrintheadRate). It also generates the margin zero data and aligns the dot data generation to the synchronization pulse from the PHI controller.

The data generator controller waits in Reset state until it receives a line start pulse from the PHI controller (data_st signal). Once a start pulse is received it proceeds to the SrclkPre state loading a counter with the SrclkPre value. While in this state it decrements the counter. No data is read or output at this stage. When the count is zero the machine proceeds to the DataGen1 state.

On transition it loads the counter with the printhead size (print_head_size). If margining is to be used then the configured print head_size should be adjusted by the dot margin value i.e. print_head_size=(physical_print_head_size−(dot_margin*2)).

Dot data is transferred to the printhead serializer in dot-pairs, with one dot-pair transferred every 3 pclk cycles. To construct a dot data pair the state machine reads one dot in the DataGen1 state, one dot in the DataGen2 state and waits for one clock cycle in the DataGen3 while the data is transferred to the data serializer. The counter will decrement for every dot data word transferred. The exact data rate is dictated by the dot buffer fill levels and the configured printhead rate (PrintheadRate). When in DataGen3 state the machine determines if it should waits for 3 cycles or transfer another dot pair to the data serializer. The generator determines the rate by comparing the rate counter (rate_cnt) with the configured PrintheadRate value. If the bit selected by the rate_cnt in the print_head_rate bus is one data is transferred, otherwise the 3 cycles are skipped (Wait1, Wait2 and Wait3). If the PrintheadRate is set to all zeros then no data will ever get transferred. The rate counter is decremented (rate_cnt) while in the DataGen2 and Wait2 states. The rate counter is allowed to wrap normally.

The pseudo-code for the rate control DataGen3 (or Wait3) state is given by:

// decrement the rate count rate_cnt −− // happens in DataGen2, or Wait2 // determine if data should be read // first determine if data is available in buffer if (rd_adr != wr_adr ) then if (print_head_rate[rate_cnt] == 1 ) then dot_active = 1 gate_srclk = 1 count −− next_state = DataGen1 else dot_active = 0 gate_srclk = 0 next_state = Wait1 else dot_active = 0 gate_srclk = 0 next_state = Wait1

When the dot counter reaches zero the state machine will jump to the MarginGen1 state if the configured margin value is non-zero, otherwise it will jump directly to the SrclkPost state. On transition to MarginGen1 state it loads the cycle counter with the dot_margin value, and begins to count down. While in the MarginGen1, MarginGen2 and MarginGen3 state machine loop the data generator logic block writes dot data to the printhead but does not read from the dot buffers. It creates zero dot data words for the margin duration. As with normal dot data, it creates one dot in MarginGen1 and MarginGen2 states, then wait a clock cycle to allow the transfer to the data serializer to complete.

When the counter reaches zero the machine jumps to the SrclkPost state, loads the clock counter with the SrclkPost value and decrements. When the count is finished the state machine returns to the Reset and awaits the next start pulse. Should a line sync arrive before the data generators have completed (data_fin signal) the PHI controller will detect a print error and stall the PHI interface.

As a consequence of the data transfer mechanism of dot pair cycles followed by a wait state, the printhead size (print_head_size) and dot margin (dot_margin) must always be even dot values.

32.9.10.3 Data Serializer

The data serializer block converts 12-bit dot data at pclk rates (nominally 160 MHz) to 2-bit data at doclk rates (nominally 320 MHz).

The srclk is only active when data is available for transfer to the printhead, as enabled by the gate_srclk signal. The data rate mechanism in the data generator block will mean that data is not transferred to the printhead on every set of 3 pclk cycles. Both the dot_data and gate_srclk signals are controlled by the data generator block and can only change on a fixed 3 pclk cycle boundary. Data is transferred to the printhead on both edges of srclk (i.e double data rate DDR). Directly after a line sync pulse the mux control logic and the srclk generation logic are reset to a known state (the srclk is set high). Before data can begin transfer to the printhead it must generate a line setup edge on srclk, causing srclk to go low. The line setup edge happens SrclkPre number of pclk cycles after the line sync falling edge (indicated by the sr_init signal from the data generator block).

All data transfers to the printhead will be in groups of 6 2-bit data words, each word clocked on an edge of srclk. For each group srclk will start low and end low.

At the end of a full line of data transfer the srclk must generate a line complete edge to return the srclk to a high state before the next line sync pulse. The data generator block generates a sr_com signal to indicate that the data transfer to the printhead has completed and that the line complete edge can be inserted. The sr_com signal is generated before the SrClkPost period.

The data serializer block allows easy separation of clock gating and clock to logic structures from the rest of the PHI interface.

The mux logic determines which data bits from the dot_data bus should be selected for output on the ph_data bus to the printhead. The mux selector is initialized by an edge detect on the sr_init signal from the data generator.

// determine wrap and init points if (phi_serial_order == 1) then mux_wrap = 5 mux_init = 0 else mux_wrap = 0 mux_init = 5 // the mux selector logic if ((sr_init_edge == 1)OR( mux_sel == mux_wrap )) then mux_sel = mux_init elsif ( phi_serial_order == 1 ) then mux_sel−− // decrement order else mux_sel++ // increment order

The dot data serialization order can be configured by PhiSerialOrder register. If the PhiSerialOrder is zero the order is dot[1:0], then dot[3:2] then dot[5:4]. If the register is one then the order is dot[5:4], dot[3:2], dot[1:0].

The srclk control logic is initialized to 1 when a line_st positive edge is detected. If either sr_com_edge, sr_init_edge or gate_srclk are equal to one srclk is transitioned. srclk is always clocked out to the output pins on the negative edge of doclk to place the clock edge in the centre of the data.

The pseudo code for the control logic is:

if (line_st_edge ==1 ) then srclk_gen = 1 elsif ((gate_srclk ==1) OR (sr_init_edge==1) OR (sr_com_edge==1)) then srclk_gen = ~srclk_gen else // hold 33 Package and Test Test Units 33.1 JTAG Interface

A standard JTAG (Joint Test Action Group) Interface is included in SoPEC for Bonding and 10 testing purposes. The JTAG port will provide access to all internal BIST (Built In Self Test) structures.

33.2 Scan Test I/O

The SoPEC device will require several test IO's for running scan tests. In general scan in and scan out pins will be multiplexed with functional pins.

33.3 Analog Test Units

33.3.1 USB PHY Testing

The USB phy analog macro, will contain built-in in test structure, which can be access by either the CPU or through the JTAG port.

33.3.2 Embedded PLL Testing

The embedded clock generator PLL will require test access from JTAG port.

34 SoPEC Pinning and Package

34.1 Overview

It is intended that the SoPEC package be a 100 pin LQFP. Any spare pins in the package may be used by increasing the number of available GPIO pins or adding extra power and ground pin. The pin list shows the minimum pin requirement for the SoPEC device.

TABLE 217 SoPEC Pin List (100 LQFP) I/O Test Rate Freq IO Cell Test Macro Group Pin Name #pins Dir Type Volt (S/D) (Mhz) Description Type Function Function Clocks and resets Group 1 Xtalin 1 I N/A N/A 32 Crystal AINSA_(—) None Input pin PM_A Xtalout 1 O N/A N/A 32 Crystal ABNST_(—) None output pin PM_A Group 2 reset_n 1 I LVTTL 3.3 v s 10 Asynchronous IT33LTPUT_(—) LT active low PM_A (leakage reset test) PrintHead Interface Group 3 phead_(—) 8 O LVDS 1.5 v d 160 Print head OLVDS15_(—) None data data PM_A Srclk 4 O LVDS 1.5 v d 160 Print head OLVDS15_(—) None clock PM_A Group 4 Readl 1 O LVTTL 3.3 v s 160 Common BT3365T_(—) A_Clock Print head PM_A mode control Frclk 1 I/O LVTTL 3.3 v s 160 Common BT3365T_(—) B_Clock Fire pattern PM_A shift clock, needs to toggle once per fire cycle phi_spare 1 I/O LVTTL 3.3 v s 160 PHI spare BT3365T_(—) C_Clock1 pin (old PM_A profile pin) Lsyncl 1 I/O LVTTL 3.3 v s 160 Line Sync BT3365T_(—) C_Clock2 output from PM_A Master to Slaves USB Connections Group 5 Usb_hostd 2 I/O Differ- 3.3 v s 12 USB BUSB2_(—) None ential differential PM_A data for host Usb_devd 2 I/O Differ- 3.3 v s 12 USB BUSB2_(—) None ential differential PM_A data for device Group 6 usbd_(—) 1 I LVTTL 3.3 v s 10 USB device BT3365T_(—) 1 scan out vbus_(—) VBUS PM_C sense power sense usbd_(—) 1 O LVTTL 3.3 v s 10 USB device BT3365T_(—) 1 scan out pull_(—) termination PM_C up_en enable JTAG Group 7 Tdo 1 O LVTTL 3.3 v s 10 JTAG Test BT3365T_(—) C_Clock3 data out PM_A port Tms 1 I LVTTL 3.3 v s 10 JTAG Test IT33RIT_(—) RI mode select PM_A Tdi 1 I LVTTL 3.3 v s 10 JTAG Test IT33D1PUT_(—) DI1 data in port PM_A Tck 1 I LVTTL 3.3 v s 10 JTAG Test IT33D2PUT_(—) DI2 access port PM_A clock General Purpose IO Group 8 Gpio 4 I/O LVTTL 3.3 v s 32 ISI BT3335PUT_(—) 4 Scanin [3:0] interface PM_B pins/GPIO Group 9 Gpio 4 I/O High 3.3 v s 32 LED driver BT3365T_(—) 4 Scanin PCNT [7:4] Drive pins/ PM_C PROGSROM LVTTL general OSC purpose Input/Output Group Gpio 12 I/O LVTTL 3.3 v s 32 General BT3365PUT_(—) 2 Scanin DIAGOUT 10 [19:8] purpose PM_B 10 Scanout (aka Input/Output MRSTR0) Group Gpio 3 I/O LVTTL 3.3 v s 32 General BT3365PUT_(—) CE0_Scan 11 [22:20] purpose PM_B TESTM3 Input/Output TSTN1 Group Gpio 10 I/O LVTTL 3.3 v s 32 Functional BT3365T_(—) 6 Scanin 12 [31:23] Spare IOs PM_C 4 scanout required for scan test Analog Power IO Group agnd 1 I Power N/A N/A N/A PLL analog AINSD3_(—) None 13 gnd PM_A avdd 1 I Power N/A N/A N/A PLL analog AINSD3_(—) None vdd PM_A agnd 1 I Power N/A N/A N/A Oscillator AINSD_(—) None analog gnd PM_A avdd 1 I Power N/A N/A N/A Oscillator AINSD_(—) None analog vdd PM_A Test Only Pin Group TE 1 I CMOS 1.5 v N/A N/A Test Enable IC15TEPDT_(—) Test only 14 PM_A VPP 1 I CMOS 1.5 v N/A N/A Fat Wire DRAMVPP_(—) Test only Analog PM Receiver/D river for Embedded DRAM Analog Inputs VWP 1 I CMOS 1.5 v N/A N/A Fat Wire DRAMVWP_(—) Test only Analog PM Receiver/D river for Embedded DRAM Analog Inputs VREFX 1 I CMOS 1.5 v N/A N/A Fat Wire DRAMVREFX_(—) Test only Analog PM Receiver/D river for Embedded DRAM Analog Inputs DLT 1 I CMOS 1.5 v N/A N/A DRAM IC15DLTPUT_(—) Test only Iddq Test PM MC 1 I CMOS 1.5 v N/A N/A IO Mode IC15MCT_(—) Test only Control PM_A DRAM_(—) 1 I CMOS 1.5 v N/A N/A DRAM IC15LTPUT_(—) Test only EN Enable(EN) PM_A Total Signal Pins 73 Functional pin count is 62 Test IO count 51 Power Only Pins Group Gnd 8 I Power N/A N/A N/A gnd GND_(—) None 15 PM_A Vdd 4 I Power N/A N/A N/A vdd 1.5 v, VDD150_(—) None core PM_A voltage vdd330 4 I Power N/A N/A N/A vdd 3.3 v, VDD330_(—) None IO voltage PM_A Group vdd/gnd 11 I Power N/A N/A N/A Power pin GND_PM_A/ None 15 fill, VDD150_(—) GND.Vdd1.5, PM_A/ Vdd3.3 VDD330_(—) as required PM_A Total Pins 100 Bilithic Printheads 1 Background

Silverbrook's bilithic Memjet™ printheads are the target printheads for printing systems which will be controlled by SoPEC and MoPEC devices.

This document presents the format and structure of these printheads, and describes the their possible arrangements in the target systems. It also defines a set of terms used to differentiate between the types of printheads and the systems which use them.

Bilithic Printhead Configurations

2 Definitions

This document presents terminology and definitions used to describe the bilithic printhead systems. These terms and definitions are as follows:

-   -   Printhead Type—There are 3 parameters which define the type of         printhead used in a system:     -   Direction of the data flow through the printhead (clockwise or         anti-clockwise, with the printhead shooting ink down onto the         page).     -   Location of the left-most dot (upper row or lower row, with         respect to V₊,).     -   Printhead footprint (type A or type B, characterized by the data         pin being on the left or the right of V₊, where V₊ is at the top         of the printhead).     -   Printhead Arrangement—Even though there are 8 printhead types,         each arrangement has to use a specific pairing of printheads, as         discussed in Section 3. This gives 4 pairs of printheads.         However, because the paper can flow in either direction with         respect to the printheads, there are a total of eight possible         arrangements, e.g. Arrangement 1 has a Type 0 printhead on the         left with respect to the paper flow, and a Type 1 printhead on         the right. Arrangement 2 uses the same printhead pair as         Arrangement 1, but the paper flows in the opposite direction.     -   Color 0 is always the first color plane encountered by the         paper.     -   Dot 0 is defined as the nozzle which can print a dot in the         left-most side of the page.     -   The Even Plane of a color corresponds to the row of nozzles that         prints dot 0.

Note that in all of the relevant drawings, printheads should be interpreted as shooting ink down onto the page.

FIG. 295 shows the 8 different possible printhead types. Type 0 is identical to the Right Printhead presented in FIG. 297 in [1], and Type 1 is the same as the Left Printhead as defined in [1].

While the printheads shown in FIG. 295 look to be of equal width (having the same number of nozzles) it is important to remember that in a typical system, a pair of unequal sized printheads may be used.

2.1 Combining Bilithic Printheads

Although the printheads can be physically joined in the manner shown in FIG. 296, it is preferable to provide an arrangment that allows greater spacing between the 2 printheads will be required for two main reasons:

-   -   inaccuracies in the backetch     -   cheaper manufacturing cost due to decreasing the tolerance         requirements in sealing the ink reservoirs behind the printhead

Failing to account for these inaccuracies and tolerances can lead to misalignment of the nozzle rows both vertically and horizontally, as shown in FIG. 297.

An even row of color n on printhead A may be vertically misaligned from the even row of color n on printhead B by some number of dots e.g. in FIG. 297 this is shown to be 5 dots. And there can also be horizontal misalignment, in that the even row of color n printhead A is not necessarily aligned with the even row of color n+1 on printhead A, e.g. in FIG. 297 this horizontal misalignment is 6 dots.

The resultant conceptual printhead definition, shown in FIG. 297 has properties that are appropriately parameterized in SoPEC and MoPEC to cater for this class of printheads.

The preferred printheads can be characterized by the following features:

-   -   All nozzle rows are the same length (although may be         horizontally displaced some number of dots even within a color         on a single printhead)     -   The nozzles for color n printhead A may not be printing on the         same line of the page as the nozzles for color n printhead B. In         the example shown in FIG. 297, there is a 5 dot displacement         between adjacent rows of the printheads.     -   The exact shape of the join is an arbitrary shape although is         most likely to be sloping (if sloping, it could be sloping         either direction)     -   The maximum slope is 2 dots per row of nozzles     -   Although shift registers are provided in the printhead at the 2         sides of the joined printhead, they do not drive nozzles—this         means the printable area is less than the actual shift         registers, as highlighted by FIG. 298.         2.2 Printhead Arrangements

Table 218 defines the printhead pairing and location of the each printhead type, with respect to the flow of paper, for the 8 possible arrangements

Printhead on left side, Printhead on right side, Printhead with respect to the flow with respect to the flow of Arrangement of paper paper Arrangement 1 Type 0 Type 1 Arrangement 2 Type 1 Type 0 Arrangement 3 Type 2 Type 3 Arrangement 4 Type 3 Type 2 Arrangement 5 Type 4 Type 5 Arrangement 6 Type 5 Type 4 Arrangement 7 Type 6 Type 7 Arrangement 8 Type 7 Type 6 3 Bilithic Printhead Systems

When using the bilithic printheads, the position of the power/gnd bars coupled with the physical footprint of the printheads mean that we must use a specific pairing of printheads together for printing on the same side of an A4 (or wider) page, e.g. we must always use a Type 0 printhead with a Type 1 printhead etc.

While a given printing system can use any one of the eight possible arrangements of printheads, this document only presents two of them, Arrangement 1 and Arrangement 2, for purposes of illustration. These two arrangements are discussed in subsequent sections of this document. However, the other 6 possibilities also need to be considered.

The main difference between the two printhead arrangements discussed in this document is the direction of the paper flow. Because of this, the dot data has to be loaded differently in Arrangement 1 compared to Arrangement 2, in order to render the page correctly.

3.1 Example 1

Printhead Arrangement 1

FIG. 299 shows an Arrangement 1 printing setup, where the bilithic printheads are arranged as follows:

-   -   The Type 0 printhead is on the left with respect to the         direction of the paper flow.     -   The Type 1 printhead is on the right.

Table 219 lists the order in which the dot data needs to be loaded into the above printhead system, to ensure color 0-dot 0 appears on the left side of the printed page.

TABLE 219 Order in which the even and odd dots are loaded for printhead Arrangement 1 Type 0 printhead Type 1 printhead Dot Sense when on the left when on the right Odd Loaded second in Loaded first in descending order. descending order. Even Loaded first in Loaded second in ascending order. ascending order.

FIG. 300 shows how the dot data is demultiplexed within the printheads.

FIG. 301 and FIG. 302 show the way in which the dot data needs to be loaded into the printheads in Arrangement 1, to ensure that color 0-dot 0 appears on the left side of the printed page. Note that no data is transferred to the printheads on the first and last edges of SrClk.

3.2 Example 2

Printhead Arrangement 2

FIG. 303 shows an Arrangement 2 printing setup, where the bilithic printheads are arranged as follows:

-   -   The Type 1 printhead is on the left with respect to the         direction of the paper flow.     -   The Type 0 printhead is on the right.

Table 220 lists the order in which the dot data needs to be loaded into the above printhead system, to ensure color 0-dot 0 appears on the left side of the printed page.

TABLE 220 Order in which the even and odd dots are loaded for printhead Arrangement 2 Type 0 printhead Type 1 printhead Dot Sense when on the right when on the left Odd Loaded first in Loaded second in descending order. descending order. Even Loaded second in Loaded first in ascending order. ascending order.

FIG. 304 shows how the dot data is demultiplexed within the printheads.

FIG. 305 and FIG. 306 show the way in which the dot data needs to be loaded into the printheads in Arrangement 2, to ensure that color 0-dot 0 appears on the left side of the printed page.

Note that no data is transferred to the printheads on the first and last edges of SrClk.

4 Conclusions

Comparing the signalling diagrams for Arrangement 1 with those shown for Arrangement 2, it can be seen that the color/dot sequence output for a printhead type in Arrangement 1 is the reverse of the sequence for same printhead in Arrangement 2 in terms of the order in which the color plane data is output, as well as whether even or odd data is output first. However, the order within a color plane remains the same, i.e. odd descending, even ascending.

From FIG. 307 and Table 221, it can be seen that the plane which has to be loaded first (i.e. even or odd) depends on the arrangement. Also, the order in which the dots have to be loaded (e.g. even ascending or descending etc.) is dependent on the arrangement.

As well as having a mechanism to cope with the shape of the join between the printheads, as discussed in Section 2.1, if the device controlling the printheads can re-order the bits according to the following criteria, then it should be able to operate in all the possible printhead arrangements:

-   -   Be able to output the even or odd plane first.     -   Be able to output even and odd planes in either ascending or         descending order, independently.     -   Be able to reverse the sequence in which the color planes of a         single dot are output to the printhead.

TABLE 221 Order in which even and odd dots and planes are loaded into the various printhead arrangements Printhead Left side of Right side of Arrangement printed page printed page Arrangement 1 Even ascending Odd descending loaded first loaded first Odd descending Even ascending loaded second loaded second Arrangement 2 Odd descending Even ascending loaded first loaded first Even ascending Odd descending loaded second loaded second Arrangement 3 Odd ascending Even descending loaded first loaded first Even descending Odd ascending loaded second loaded second Arrangement 4 Even descending Odd ascending loaded first loaded first Odd ascending Even descending loaded second loaded second Arrangement 5 Odd ascending Even descending loaded first loaded first Even descending Odd ascending loaded second loaded second Arrangement 6 Even descending Odd ascending loaded first loaded first Odd ascending Even descending loaded second loaded second Arrangement 7 Even ascending Odd descending loaded first loaded first Odd descending Even ascending loaded second loaded second Arrangement 8 Odd descending Even ascending loaded first loaded first Even ascending Odd descending loaded second loaded second CMOS Support on Bilithic Printhead 1 Basic Requirements

To create a two part printhead, of A4/Letter portrait width to print a page in 2 seconds. Matching Left/Right chips can be of different lengths to make up this length facilitating increased wafer usage. the left and right chips are to be imaged on an 8 inch wafer by “Stitching” reticle images.

The memjet nozzles have a horizontal pitch of 32 um, two rows of nozzles are used for a single colour. These rows have a horizontal offset of 16 um. This gives an effective dot pitch of 16 um, or 62.5 dots per mm, or 1587.5 dots per inch, close enough to market as 1600 dpi.

The first nozzle of the right chip should have a 32 um horizontal offset from the final nozzle of the left chip for the same color row. There is no ink nozzle overlap (of the same colour) scheme employed.

1.1 Power Supply

Vdd/Vpos and Ground supply is made through 30 um wide pads along the length of the chip using conductive adhesive to bus bar beside the chips. Vdd/Vpos is 3.3 Volts. (12V was considered for Vpos but routing of CMOS Vdd at 3.3V would be a problem over the length of the chips, but this will be revisited).

1.2 MEMS Cells

The preferred memjet device requires 180 nJ of energy to fire, with a pulse of current for 1 usec. Assuming 95% efficiency, this requires a 55 ohm actuator drawing 57.4 mA during this pulse.

1.2.1 Issue!!!

For 1 pages per 2 second, or ˜300 mm*62.5 (dots/mm)/2 sec˜=10 kHz or 100 usec per line. With 1 usec fire pulse cycle, every 100th nozzle needs to fire at the same time. We have 13824 nozzles across the page, so we fire 138 nozzles at a time.

1.2.2 64 um Unit Cell Height

This cell would have 4 line spacing between the odd and even dots, and 8 line spacing between adjacent colours.

1.2.3 80 um Unit Cell Height

This cell would have 5 line spacing between the odd and even dots, and 10 line spacing between adjacent colours.

1.3 Versions

1.3.1 6 Colour 1600 dpi with 64 um Unit Cell

Left and Right Chip.

1.3.2 6 Colour 1600 dpi with 80 um Unit Cell

Left and Right Chip.

1.3.3 4 Colour 800 dpi with 80 um Unit Cell

For camera application. Single nozzle row per colour.

1.4 Air Supply

Air must be supplied to the MEMS region through holes in the chip.

2 Head Sizes

The combined heads have 13824 nozzles per colour totalling 221.184 mm of print area. Enough to provide full breadth for A4 (210 mm) and Letter (8.5 inch or 215.9 mm).

TABLE 1 Head Combinations Left Head Right Head Stitch Nozzles Stitch Nozzles Parts per Colour Parts per Colour 8 11160 2 2664 7 9744 3 4080 6 8328 4 5496 5 6912 5 6912 4 5496 6 8328 3 4080 7 9744 2 2664 8 11160

Nozzles per Colour is calculated as ((“Stitch Parts” −1)*118+104)*12. Nozzles per row is half this value. Most likely the 8:2 head set will not be manufactured. The preferred wafer layout, manages to avoid this set, without any loses.

3 Interface

Each print head has the same I/O signals (but the Left and Right versions might have a different pin out).

TABLE 2 I/O pins Max Speed Name I/O Function Common (MHz) Data I Dot data for colours 0-5, No 320 [0-1] using Differential Signalling (DataL the complementary signal), colours[0-2] on Data[0], colour[3-5] on Data[1] DataL I complementary signal of [0-1] Data[0-1] SrClk I Dot data shift clock using No¹ 320 Differential Signalling (SrClkL the complementary signal) SrClkL I complementary signal of SrClk ReadL I FrClk, Pr, LSyncL output mode Yes 1 if signal mode bit is set FrClk I Fire pattern shift clock Yes 1 O nozzle test result Yes² (mode = 0b001), LsyncL = 0 CMOS testing (mode = 0b111), LsyncL = 1 Pr I Pulse Profile for all colours Yes 1³ O Temperature Output Yes^(b) (mode = 0b010), LsyncL = 0 CMOS testing (mode = 0b111), LsyncL = 1 LsyncL I 0 - Capture dot data for next Yes 0.1⁴ print line O CMOS testing (mode = 0b111), Yes^(b) LsyncL = 1 ¹Functionally could be common, but for timing/electrical reasons should run point to point. ²Can be shared if one side has mode = 0b000 ³1 MHz cycle, but the resolution of the mark/space ratio may require 50 ns. ⁴10 kHz cycle, with minimum low pulse of 10 ns (no maximum).

Pins marked as common can be controlled by the same signal from the controller (SOPEC).

3.1 Dot Firing

To fire a nozzle, three signals are needed. A dot data, a fire signal, and a profile signal. When all signals are high, the nozzle will fire.

The dot data is provide to the chip through a dot shift register with input Data[x], and clocked into the chip with SrClk. The dot data is multiplex on to the Data signals, as Dot[0-2] on Data[0], and Dot[3-51 on Data[2]. After the dots are shifted into the dot shift register, this data is transfer into the dot latch, with a low pulse in LsyncL. The value in the dot latch forms the dot data used to fire the nozzle. The use of the dot latch allows the next line of data to be loaded into the dot shift register, at the same time the dot pattern in the dot latch is been fired.

Across the top of a column of nozzles, containing 12 nozzles, 2 of each colour (odd and even dots, 4 or 5 lines apart), is two fire register bits and a select register bit. The fire registers forms the fire shift register that runs length of the chip and back again with one register bit in each direction flow. The select register forms the Select Shift Register that runs the length of the chip. The select register, selects which of the two fire registers is used to enables this column. A ‘0’ in this register selects the forward direction fire register, and a ‘1’ selects the reverse direction fire register. This output of this block provides the fire signal for the column.

The third signal needed, the profile, is provide for all colours with input Pr across the whole colour row at the same time (with a slight propagation delay per column).

3.2 Dot Shift Register Orientation

The left side print head (chip) and the right side print head that form complete bi-lithic print head, have different nozzle arrangement with respect to the dot order mapping of the dot shift register to the dot position on the page.

With this mapping, the following data streams will need to provided.

Left Head Right Head Size n-m dot order m 7:3 97 44 [13822,13820,13818, . . . , 40 80 [1,3,5, . . . ,4075,4077,4079,] 4084,4082,4080,] line y line y+5 [4078,4076,4074, . . . ,4,2,0] [4081,4083,4085, . . . ,13819, line y+5 13821, 13823] line y 6:4 83 28 [13822,13820,13818, . . . , 54 96 [1,3,5, . . . ,5491,5493,5495,] 5500,5498,5496,] line y line y+5 [5494,5492,5490, . . . ,4,2,0] [5497,5499,5501, . . . ,13819, line y+5 13821, 13823] line y 5:5 69 12 [13822,13820,13818, . . . ,6916, 69 12 [1,3,5, . . . ,6907,6909,6911,] 6914,6912,] line y line y+5 [6910,6908,6906, . . . ,4,2,0] [6913,6915,6917, . . . ,13819, line y+5 13821, 13823] line y 4:6 54 96 13822,13820,13818, . . . ,8332, 83 28 [1,3,5, . . . ,8323,8325,8327,] 8330,8328,] line y line y+5 [8326,8324,8322, . . . ,4,2,0] 8329,8331,8333, . . . ,13819, line y+5 13821, 13823] line y 3:7 40 80 13822,13820,13818, . . . ,9748, 97 44 1,3,5, . . . ,9739,9741,9743,] 9746,9744,] line y line y+5 9742,9740,9738, . . . ,4,2,0] 9745,97447,9749, . . . ,13819, line y+5 13821, 13823] line y

The data needs to be multiplexed onto the data pins, such that Data[0] has {(C0, C1, C2), (C0, C1, C2) . . . } in the above order, and Data[1] has {(C3, C4, C5), (C3, C4, C5) . . . }.

FIG. 311 shows the timing of data transfer during normal printing mode. Note SrClk has a default state of high and data is transferred on both edges of SrClk. If there are L nozzles per colour, SrClk would have L+2 edges, where the first and last edges do not transfer data.

Data requires a setup and hold about the both edges of SrClk. Data transfers starts on the first rising after LSyncL rising. SrClk default state is high and needs to return to high after the last data of the line. This means the first edge of SrClk (falling) after LSyncL rising, and the last edge of SrClk as it returns to the default state, no data is transferred to the print head. LSyncL rising requires setup to the first falling SrClk, and must stay high during the entire line data transfer until after last rising SrClk.

3.3 Fire Shift Register

The fire shift register controls the rate of nozzle fire. If the register is full of ‘1’s then the you could print the entire print head in a single FrClk cycle, although electrical current limitations will prevent this happening in any reasonable implementation.

Ideally, a ‘1’ is shifted in to the fire shift register, in every n^(th) position, and a ‘0’ in all other position. In this manner, after n cycles of FrClk, the entire print head will be printed.

The fire shift register and select shift registers allow the generation of a horizontal print line that on close inspection would not have a discontinuity of a “saw tooth” pattern, FIGS. 312 a) & b) but a “sharks tooth” pattern of c).

This is done by firing 2 nozzles in every 2n group of nozzle at the same time starting from the outer 2 nozzles working towards the centre two (or the starting from the centre, and working towards the outer two) at the fire rate controlled by FrClk.

To achieve this fire pattern the fire shift register and select shift register need to be set up as show in FIG. 313.

The pattern has shifted a ‘1’ into the fire shift register every n^(th) positions (where n is usually is a minimum of about 100) and n ‘1’s, followed n ‘0’s in the select shift register. At a start of a print cycle, these patterns need to be aligned as above, with the 1000 . . . ” of a forward half of fire shift register, matching an n grouping of ‘1’ or ‘0’s in the select shift register. As well, with the “1000 . . . ” of a reverse half of the fire shift register, matching an n grouping of ‘1’ or ‘0’s in the select shift register. And to continue this print pattern across the butt ends of the chips, the select shift register in each should end with a complete block of n ‘1’s (or ‘0’s).

Since the two chips can be of different lengths, initialisation of these patterns is an issue. This is solved by building initialisation circuitry into chips. This circuit is controlled by two registers, nlen(14) and count(14) and b(1). These registers are loaded serially through Data[0], while LSyncL is low, and ReadL is high with FrClk.

The scan order from input is b, n[13-0],c[0-13],color[5-0], mode[2-0] therefore b is shifted in last. The system color and mode registers are unrelated to the Fire Shift Register, but are loaded at the same time as this block. There function is described later.

TABLE 4 Head Combinations Initialisation for n = 100 count_(A) = count_(B) = Nozzle s Nozzle s nlen_((A&B)) = (L_(A)/2) mod rem = (L_(A) − L_(B) + rem) L_(B) L_(A) n − 1 n − 1 b _(A) b _(B) (L_(B)/2) mod n mod n − 1 4080 9744 99 71 0 0 40 3 5496 8328 99 63 0 0 48 79 6912 6912 99 55 0 0 56 55

The following table shows the values to programme the bi-lithic head pairs using a fire pattern length of 100. The calculation assumes head ‘A’ is the longest head of the pair and once the registers are initialised with LA FrClk cycles (ReadL=‘0’, LSyncL=‘1’). rem would be the correct value for count_(B) if chip B was only clocked (FrClk) L_(B) times. But this chip will be over clocked L_(A)−L_(B) cycles. The values of b_(A) and b_(B) are either the same or inverse of each other. The actually value does not matter. They need to be different from each other if the select shift registers would end up with different values at the butt ends. If (L_(A)/2n) is even (and count_(A) is non zero), then the final run in ‘A’s select shift register will be !b_(A). If (L_(A)−L_(B)/2) mod n is even (and count_(b) is non zero) then the final run in ‘B’s select shift register will be !b_(B).

3.4 System Registers

As describe above, the Fire Shift Register generation block, also contains some system registers.

TABLE 5 System Registers Name Size Function Color 6 Each bit is an enable for the corresponding colour. If color[X]=0, then Pr_(x) is 0 and SrClk_(x) is 0. If color[X]=1, then Pr_(x) follows the Pr signal and SrClk_(x) is deserialised SrClk. Mode 3 Mode[0] = 1, then FrClk pin is used as an output, internally the FrClk signal is set to 0 Mode[1] = 1, then Pr pin is used as an output, internally the Pr signal is set to 0 Mode[2] = 1, then LsyncL pin is used as an output, internally the LsyncL signal is set to 1 3.5 Profile Pattern

A profile pattern is repeated at FrClk rate. It is expected to be a single pulse about 1 us long. But it could be a more complicated series of pulse. The actual pattern depends on the ink type.

The following figure show the external timing to print a line of data. In this example the line is printed in 8 cycles of FrClk.

3.6 Interface Modes

The print head has eight different modes controlled by signals ReadL and LSyncL and system mode register. As seen in FIG. 318 with both LSyncL and ReadL high, the chip in normal printing mode. Some of these modes can operate at the same time, but may interfere with the result of the other modes.

TABLE 6 Print Head Modes Mode Internal ReadL LSyncL Function Register Mapping 1 1 Normal Print Mode 000 (XXX) SrClk=SrClk/3 frclk=FrClk SelClk=0 FsClk=FrClk Scan=0 CoreScan=0 X 0 Dot Load Mode 000 (XXX) Dot latches are open, loaded with Dot shift registers, latch once LSyncL returns to 1 (this happens regardless of ReadL) Enables Dot Shift register to capture fire result. 1 0 Fire Load Mode 000 (XXX) SrClk=X Data[0] will shift frclk=X through mode, color, nlen, SelClk=X count and b with FrClk FsClk=FrClk Scan=1 CoreScan=X 0 1 Reset Nozzle Test 001 SrClk=SrClk Resets the state of FrClk=FrClk nozzle test SelClk=FrClk circuit FsClk=FrClk Scan=0 CoreScan=1 0 1 CMOS testing mode 111 The contents of the dot shift registers are serial shifted out on LsyncL (colour0-1), FrClk (colour2-3), Pr (colour4-5) with SrClk 0 1 Fire Initialise mode 000 (XXO) The contents of the fire shift register and select shift register is generated with FrClk 0 0 Temperature Output 010 SrClk=X The series of Sigma frclk=0 Delta output are SelClk=0 clocked out on Pr with FsClk=0 FrClk. Scan=0 The sum of these bits CoreScan=X represent the temperature of the chip. 0 0 Nozzle Test Output 001 The result of a nozzle test is output on FrClk. 3.6.1 Printing

FIG. 318 shows show timing for normal printing. During this action, we drop out of Normal Print Mode, to Dot Load Mode between line transfers. For printing to perform correctly, all other signals should be stable.

3.6.2 Initialising for Printing

To initialise for printing the fire shift registers and select shift registers need to be setup into a state as shown in FIG. 318. To do this the chips are put into Fire Load Mode and the values for nlen, count and b are serially shifted from Data[0] clocked by FrClk. As the two chip have separate Data line, and common FrClk, this happens at the same time. Once this is done, mode is changed to Fire Initialise Mode, and further LA FrClk cycles are provided to both chips. During all these operation Pr should be low, to prevent unintentional firing for nozzles.

3.6.3 Nozzle Testing

Nozzle testing is done by firing a single nozzle at a time and monitoring the FrClk pin in the Nozzle Test Output mode.

Each nozzle has a test switch which closes when the nozzle is fired with an energy level greater than required for normal ink ejection. All 12 switches in a nozzle column are connect in parallel to the following circuit.

This circuit is initialised when ever LSyncL is high and ReadL is low (Reset Nozzle Test mode). This forces all “switch nodes” to low, and the feedback through lower NOR gate will latches this value. With LSyncL low and ReadL still low (Nozzle Test Output mode) the Testout of the first nozzle column is output on FrClk. If any switch is closed, the switch node of this column will be pulled up, and will ripple through to the output as transition from high to low.

Nozzle testing requires a setup phase in order to fire only one nozzle. There are many ways to achieve this. Simplest might be to load a single colour with 101010 through the even nozzles, and 010101 . . . for the odd nozzles (0's for all other colours), and set up a fire pattern with n=L_(A)/2. With this fire pattern only one nozzle will fire in each Pr pulse. After firing in Nozzle Test Output mode, a single FrClk will advance to next nozzle, then Reset and Test. After L_(A)/2 cycles of this testing, a single SrClk will advance the dot shift registers to setup the untested nozzles of this colour, and another L_(A)/2 cycles of FrClk, Reset and Test will finished testing this colour. Then repeat test procedure for other colours.

3.6.4 Temperature Output

This mode is not well defined yet. In this mode, Pr will output a series of ones and zeros clocked by FrClk. After a (currently unknown) number of FrClk cycles the sum of this series will represent the temperature of the chip. Clocking frequency in this mode it expected to be in the range 10 kHz-1 MHz.

The Frequency of FrClk and the number of cycles need to be programmable. Since this mode cycles FrClk, the result of fire shift register and select shift register would be changed, but in this mode FrClk is disabled to these circuit. So printing can resume without reinitialising.

3.6.5 CMOS Testing

CMOS testing is a mode meant for chip testing before MEMS as added to the chip. This mode allows the dot shift register to be shifted out on the LsyncL, FrClk and Pr pins. Much like the nozzle test mode, the nozzles are fired while LSyncL is low, but during the firing SrClk will be pulsed, loading the dot shift register with the signal that would fire the nozzle. Once captured, the result can be shifted out.

The Dot Load Mode above violates normal printing procedure by firing the nozzles (Pr) and modify the dot shift register (SrClk).

4 Reticle Layout

To make long chips we need to stitch the CMOS (and MEMS) together by overlapping the reticle stepping field. The reticle will contain two areas:

The top edge of Area 2, PAD END contains the pads that stitch on bottom edge of Area 1, CORE. Area 1 contains the core array of nozzle logic. The top edge of Area 1 will stitch to the bottom edge of itself. Finally the bottom edge of Area 2, BUTT END will stitch to the top edge of Area 1. The BUTT END to used to complete a feedback wiring and seal the chip.

The above region will then be exposed across a wafer bottom to top. Area 2, Area 1, Area 1 . . . , Area 2. Only the PAD END of Area 2 needs to fit on the wafer. The final exposure of Area 2 only requires the BUTT END on the wafer.

4.1 TSMC U-Frame Requirements.

TSMC will be building us frames 10 mm×0.23 mm which will be placed either side of both Area 1 and Area 2.

TSMC requires 6 mm area for blading between the two exposure area. This translates to 3 mm on the reticle, as some reticules are 2× size, while most are 5×, the worst case must be used.

Security Overview

1 Introduction

A number of hardware, software and protocol solutions to security issues have been developed. These range from authorization and encryption protocols for enabling secure communication between hardware and software modules, to physical and electrical systems that protect the integrity of integrated circuits and other hardware.

It should be understood that in many cases, principles described with reference to hardware such as integrated circuits (ie, chips) can be implemented wholly or partly in software running on, for example, a computer. Mixed systems in which software and hardware (and combinations) embody various entities, modules and units can also be constructed using may of these principles, particularly in relation to authorization and authentication protocols. The particular extent to which the principles described below can be translated to or from hardware or software will be apparent to one skilled in the art, and so will not always explicitly be explained.

It should also be understood that many of the techniques disclosed below have application to many fields other than printing. Some specific examples are described towards the end of this description.

A “QA Chip” is a quality assurance chip can allows certain security functions and protocols to be implemented. The preferred QA Chip is described in some detail later in this specification.

1.5 QA Chip Terminology

The Authentication Protocols documents [5] and [6] refer to QA Chips by their function in particular protocols:

-   -   For authenticated reads in [5], ChipR is the QA Chip being read         from, and ChipT is the QA Chip that identifies whether the data         read from ChipR can be trusted. ChipR and ChipT are referred to         as Untrusted QA Device and Trusted QA Device respectively in         [6].     -   For replacement of keys in [5], ChipP is the QA Chip being         programmed with the new key, and ChipF is the factory QA Chip         that generates the message to program the new key. ChipF is         referred to as the Key Programmer QA Device in [6].     -   For upgrades of data in memory vectors in [5], ChipU is the QA         Chip being upgraded, and ChipS is the QA Chip that signs the         upgrade value. ChipS is referred to as the Value Upgrader QA         Device and Parameter Upgrader QA Device in [6].

Any given physical QA Chip will contain functionality that allows it to operate as an entity in some number of these protocols.

Therefore, wherever the terms ChipR, ChipT, ChipP, ChipF, ChipU and ChipS are used in this document, they are referring to logical entities involved in an authentication protocol as defined in [5] and [6].

Physical QA Chips are referred to by their location. For example, each ink cartridge may contain a QA Chip referred to as an INK_QA, with all INK_QA chips being on the same physical bus. In the same way, the QA Chip inside the printer is referred to as PRINTER_QA, and will be on a separate bus to the INK_QA chips.

2 Requirements

2.1 Security

When applied to a printing environment, the functional security requirements for the preferred embodiment are:

-   -   Code of QA chip owner or licensee co-existing safely with code         of authorized OEMs     -   Chip owner/licensee operating parameters authentication     -   Parameters authentication for authorized OEMs     -   Ink usage authentication

Each of these is outlined in subsequent sections.

The authentication requirements imply that:

-   -   OEMs and end-users must not be able to replace or tamper with QA         chip manufacturer/owner's program code or data     -   OEMs and end-users must not be able to perform unauthorized         activities for example by calling chip manufacturer/owner's code     -   End-users must not be able to replace or tamper with OEM program         code or data     -   End-users must not be able to call unauthorized functions within         OEM program code     -   Manufacturer/owner's development program code must not be         capable of running on all SoPECs.     -   OEMs must be able to test products at their highest upgradable         status, yet not be able to ship them outside the terms of their         license     -   OEMs and end-users must not be able to directly access the print         engine pipeline (PEP) hardware, the LSS Master (for QA Chip         access) or any other peripheral block with the exception of         operating system permitted GPIO pins and timers.         2.1.1 QA Manufacturer/Owner Code and OEM Program Code         Co-Existing Safely

SoPEC includes a CPU that must run both manufacturer/owner program code and OEM program code. The execution model envisaged for SoPEC is one where Manufacturer/owner program code forms an operating system (O/S), providing services such as controlling the print engine pipeline, interfaces to communications channels etc. The OEM program code must run in a form of user mode, protected from harming the Manufacturer/owner program code. The OEM program code is permitted to obtain services by calling functions in the O/S, and the O/S may also call OEM code at specific times. For example, the OEM program code may request that the O/S call an OEM interrupt service routine when a particular GPIO pin is activated.

In addition, we may wish to permit the OEM code to directly call functions in Manufacturer/owner code with the same permissions as the OEM code. For example, the Manufacturer/owner code may provide SHA1 as a service, and the OEM could call the SHA1 function, but execute that function with OEM permissions and not Silverbook permissions.

A basic requirement then, for SoPEC, is a form of protection management, whereby Manufacturer/owner and OEM program code can co-exist without the OEM program code damaging operations or services provided by the Manufacturer/owner O/S. Since services rely on SoPEC peripherals (such as USB2 Host, LSS Master, Timers etc) access to these peripherals should also be restricted to Manufacturer/owner program code only.

2.1.2 Manufacturer/Owner Operating Parameters Authentication

A particular OEM will be licensed to run a Print Engine with a particular set of operating parameters (such as print speed or quality). The OEM and/or end-user can upgrade the operating license for a fee and thereby obtain an upgraded set of operating parameters.

Neither the OEM nor end-user should be able to upgrade the operating parameters without paying the appropriate fee to upgrade the license. Similarly, neither the OEM nor end-user should be able to bypass the authentication mechanism via any program code on SoPEC. This implies that OEMs and end-users must not be able to tamper with or replace Manufacturer/owner program code or data, nor be able to call unauthorized functions within Manufacturer/owner program code.

However, the OEM must be capable of assembly-line testing the Print Engine at the upgraded status before selling the Print Engine to the end-user.

2.1.3 OEM Operating Parameters Authentication

The OEM may provide operating parameters to the end-user independent of the Manufacturer/owner operating parameters. For example, the OEM may want to sell a franking machine¹. ¹a franking machine print stamps

The end-user should not be able to upgrade the operating parameters without paying the appropriate fee to the OEM. Similarly, the end-user should not be able to bypass the authentication mechanism via any program code on SoPEC. This implies that end-users must not be able to tamper with or replace OEM program code or data, as well as not be able to tamper with the PEP blocks or service-related peripherals.

2.2 Acceptable Compromises

If an end user takes the time and energy to hack the print engine and thereby succeeds in upgrading the single print engine only, yet not be able to use the same keys etc on another print engine, that is an acceptable security compromise. However it doesn't mean we have to make it totally simple or cheap for the end-user to accomplish this.

Software-only attacks are the most dangerous, since they can be transmitted via the internet and have no perceived cost. Physical modification attacks are far less problematic, since most printer users are not likely to want their print engine to be physically modified. This is even more true if the cost of the physical modification is likely to exceed the price of a legitimate upgrade.

2.3 Implementation Constraints

Any solution to the requirements detailed in Section 2.1 should also meet certain preferred implementation constraints. These are:

-   -   No flash memory inside SoPEC     -   SoPEC must be simple to verify     -   Manufacturer/owner program code must be updateable     -   OEM program code must be updateable     -   Must be bootable from activity on USB2     -   Must be bootable from an external ROM to allow stand-alone         printer operation     -   No extra pins for assigning IDs to slave SoPECs     -   Cannot trust the comms channel to the QA Chip in the printer         (PRINTER_QA)     -   Cannot trust the comms channel to the QA Chip in the ink         cartridges (INK_QA)     -   Cannot trust the USB comms channel

These constraints are detailed below.

2.3.1 No Flash Memory Inside SoPEC

The preferred embodiment of SoPEC is intended to be implemented in 0.13 micron or smaller. Flash memory will not be available in any of the target processes being considered.

2.3.2 SoPEC Must be Simple to Verify

All combinatorial logic and embedded program code within SoPEC must be verified before manufacture. Every increase in complexity in either of these increases verification effort and increases risk.

2.3.3 Manufacturer/Owner Program Code Must be Updateable

It is neither possible nor desirable to write a single complete operating system that is:

-   -   verified completely (see Section 2.3.1)     -   correct for all possible future uses of SoPEC systems     -   finished in time for SoPEC manufacture

Therefore the complete Manufacturer/owner program code must not permanently reside on SoPEC. It must be possible to update the Manufacturer/owner program code as enhancements to functionality are made and bug fixes are applied.

In the worst case, only new printers would receive the new functionality or bug fixes. In the best case, existing SoPEC users can download new embedded code to enable functionality or bug fixes. Ideally, these same users would be obtaining these updates from the OEM website or equivalent, and not require any interaction with Manufacturer/owner.

2.3.4 OEM Program Code Must be Updateable

Given that each OEM will be writing specific program code for printers that have not yet been conceived, it is impossible for all OEM program code to be embedded in SoPEC at the ASIC manufacture stage.

Since flash memory is not available (see Section 2.3.1), OEMs cannot store their program code in on-chip flash. While it is theoretically possible to store OEM program code in ROM on SoPEC, this would entail OEM-specific ASICs which would be prohibitively expensive. Therefore OEM program code cannot permanently reside on SoPEC.

Since OEM program code must be downloadable for SoPEC to execute, it should therefore be possible to update the OEM program code as enhancements to functionality are made and bug fixes are applied.

In the worst case, only new printers would receive the new functionality or bug fixes. In the best case, existing SoPEC users can download new embedded code to enable functionality or bug fixes. Ideally, these same users would be obtaining these updates from the OEM website or equivalent, and not require any interaction with Manufacturer/owner.

2.3.5 Must be Bootable From Activity on USB2

SoPEC can be placed in sleep mode to save power when printing is not required. RAM is not preserved in sleep mode. Therefore any program code and data in RAM will be lost. However, SoPEC must be capable of being woken up by the host when it is time to print again.

In the case of a single SoPEC system, the host communicates with SoPEC via USB2. From SoPEC's point of view, it is activity on the USB2 device port that signals the time to wake up.

In the case of a multi-SoPEC system, the host typically communicates with the Master SoPEC chip (as above), and then the Master relays messages to other Slave SoPECs by sending data out USB2 host port(s) and into the Slave SoPEC's device port. The net result is that the Slave SoPECs and the Master SoPEC all boot as a result of activity on the USB2 device port.

Therefore SoPEC must be capable of being woken up by activity on the USB2 device port.

2.3.6 Must be Bootable From an External ROM to Allow Stand-Alone Printer Operation

SoPEC must also support the case where the printer is not connected to a PC (or the PC is currently turned off), and a digital camera or equivalent is plugged into the SoPEC-based printer. In this case, the entire printing application needs to be present within the hardware of the printer. Since the Manufacturer/owner program code and OEM program code will vary depending on the application (see Section 2.3.3 and Section 2.3.4), it is not possible to store the program in SoPEC's ROM.

Therefore SoPEC requires a means of booting from a non-PC host. It is possible that this could be accomplished by the OEM adding a USB2-host chip to the printer and simulating the effect of a PC, and thereby download the program code. This solution requires the boot operation to be based on USB2 activity (see Section 2.3.5). However this is an unattractive solution since it adds microprocessor complexity and component cost when only a ROM-equivalent was desired.

As a result SoPEC should ideally be able to boot from an external ROM of some kind. Note that booting from an external ROM means first booting from the internal ROM, and then downloading and authenticating the startup section of the program from the external ROM. This is not the same as simply running program code in-situ within an external ROM, since one of the security requirements was that OEMs and end-users must not be able to replace or tamper with Manufacturer/owner program code or data, i.e. we never want to blindly run code from an external ROM.

As an additional point, if SoPEC is in sleep mode, SoPEC must be capable of instigating the boot process due to activity on a programmable GPIO. e.g. a wake-up button. This would begin addition to the standard power-on booting.

2.3.7 No Extra Pins to Assign IDs to Slave SoPECs

In a single SoPEC system the host only sends data to the single SoPEC. However in a multi-SoPEC system, each of the slaves needs to be uniquely identifiable in order to be able for the host to send data to the correct slave.

Since there is no flash on board SoPEC (Section 2.3.1) we are unable to store a slave ID in each SoPEC. Moreover, any ROM in each SoPEC will be identical.

It is possible to assign n pins to allow 2^(n) combinations of IDs for slave SoPECs. However a design goal of SoPEC is to minimize pins for cost reasons, and this is particularly true of features only used in multi-SoPEC systems.

The design constraint requirement is therefore to allow slaves to be IDed via a method that does not require any extra pins. This implies that whatever boot mechanism that satisfies the security requirements of Section 2.1 must also be able to assign IDs to slave SoPECs.

2.3.8 Cannot Trust the Comms Channel to the QA Chip in the Printer (PRINTER_QA)

If the printer operating parameters are stored in the non-volatile memory of the Print Engine's on-board PRINTER_QA chip, both Manufacturer/owner and OEM program code cannot rely on the communication channel being secure. It is possible for an attacker to eavesdrop on communications to the PRINTER_QA chip, replace the PRINTER_QA chip and/or subvert the communications channel. It is also possible for this to be true during manufacture of the circuit board containing the SoPEC and the PRINTER_QA chip.

2.3.9 Cannot Trust the Comms Channel to the QA Chip in the Ink Cartridges (INK_QA)

The amount of ink remaining for a given ink cartridge is stored in the non-volatile memory of that ink cartridge's INK_QA chip. Both Manufacturer/owner and OEM program code cannot rely on the communication channel to the INK_QA being secure. It is possible for an attacker to eavesdrop on communications to the INK_QA chip, to replace the INK_QA chip and/or to subvert the communications channel. It is also possible for this to be true during manufacture of the consumable containing the INK_QA chip.

2.3.10 Cannot Trust the Inter-SoPEC Comms Channel (USB2)

In a multi-SoPEC system, or in a single-SoPEC system that has a non-USB2 connection to the host, a given SoPEC will receive its data over a USB2 host port. It is quite possible for an end-user to insert a chip that eavesdrops on and/or subverts the communications channel (for example performs man-in-the-middle attacks).

3 Proposed Solution

A proposed solution to the requirements of Section 2, can be summarised as:

-   -   Each SoPEC has a unique id     -   CPU with user/supervisor mode     -   Memory Management Unit     -   The unique id is not cached     -   Specific entry points in O/S     -   Boot procedure, including authentication of program code and         operating parameters     -   SoPEC physical identification         3.1 Each SoPEC Has a Unique ID

Each SoPEC needs to contains a unique SoPEC_id of minimum size 64-bits. This SoPEC_id is used to form a symmetric key unique to each SoPEC: SoPEC_id_key. On SoPEC we make use of an additional 112-bit ECID² macro that has been programmed with a random number on a per-chip basis. Thus SoPEC_id is the 112-bit macro, and the SoPEC_id_key is a 160-bit result obtained by SHA1(SoPEC_ID). ²Electronic Chip Id

The verification of operating parameters and ink usage depends on SoPEC_id being difficult to determine. Difficult to determine means that someone should not be able to determine the id via software, or by viewing the communications between chips on the board. If the SoPEC_id is available through running a test procedure on specific test pins on the chip, then depending on the ease by which this can be done, it is likely to be acceptable.

It is important to note that in the proposed solution, compromise of the SoPEC_id leads only to compromise of the operating parameters and ink usage on this particular SoPEC. It does not compromise any other SoPEC or all inks or operating parameters in general.

It is ideal that the SoPEC_id be random, although this is unlikely to occur on standard manufacture processes for ASICs. If the id is within a small range however, it will be able to be broken by brute force. This is why 32-bits is not sufficient protection.

3.2 CPU with User/Supervisor Mode

SoPEC contains a CPU with direct hardware support for user and supervisor modes. At present, the intended CPU is the LEON (a 32-bit processor with an instruction set according to the IEEE-1754 standard. The IEEE1754 standard is compatible with the SPARC V8 instruction set).

Manufacturer/owner (operating system) program code will run in supervisor mode, and all OEM program code will run in user mode.

3.3 Memory Management Unit

SoPEC contains a Memory Management Unit (MMU) that limits access to regions of DRAM by defining read, write and execute access permissions for supervisor and user mode. Program code running in user mode is subject to user mode permission settings, and program code running in supervisor mode is subject to supervisor mode settings.

A setting of 1 for a permission bit means that type of access (e.g. read, write, execute) is permitted. A setting of 0 for a read permission bit means that that type of access is not permitted.

At reset and whenever SoPEC wakes up, the settings for all the permission bits are 1 for all supervisor mode accesses, and 0 for all user mode accesses. This means that supervisor mode program code must explicitly set user mode access to be permitted on a section of DRAM.

Access permission to all the non-valid address space should be trapped, regardless of user or supervisor mode, and regardless of the access being read, execute, or write.

Access permission to all of the valid non-DRAM address space (for example the PEP blocks) is supervisor read/write access only (no supervisor execute access, and user mode has no acccess at all) with the exception that certain GPIO and Timer registers can also be accessed by user code. These registers will require bitwise access permissions. Each peripheral block will determine how the access is restricted.

With respect to the DRAM and PEP subsystems of SoPEC, typically we would set user read/write/execute mode permissions to be 1/1/0 only in the region of memory that is used for OEM program data, 1/0/1 for regions of OEM program code, and 0/0/0 elsewhere (including the trap table). By contrast we would typically set supervisor mode read/write/execute permissions for this memory to be 1/1/0 (to avoid accidentally executing user code in supervisor mode).

The SoPEC_id parameter (see Section 3.1) should only be accessible in supervisor mode, and should only be stored and manipulated in a region of memory that has no user mode access. μ3.4 Unique ID IS Not Cached

The unique SoPEC_id needs to be available to supervisor code and not available to user code. This is taken care of by the MMU (Section 3.3).

However the SoPEC_id must also not be accessable via the CPU's data cache or register windows. For example, if the user were to cause an interrupt to occur at a particular point in the program execution when the SoPEC_id was being manipulated, it must not be possible for the user program code to turn caching off and then access the SoPEC_id inside the data cache. This would bypass any MMU security.

The same must be true of register windows. It must not be possible for user mode program code to read or modify register settings in a supervisor program's register windows.

This means that at the least, the SoPEC_id itself must not be cacheable. Likewise, any processed form of the SoPEC_id such as the SoPEC_id key (e.g. read into registers or calculated expected results from a QA_Chip) should not be accessable by user program code.

3.5 Specific Entry Points in O/S

Given that user mode program code cannot even call functions in supervisor code space, the question arises as how OEM programs can access functions, or request services. The implementation for this depends on the CPU.

On the LEON processor, the TRAP instruction allows programs to switch between user and supervisor mode in a controlled way. The TRAP switches between user and supervisor register sets, and calls a specific entry point in the supervisor code space in supervisor mode. The TRAP handler dispatches the service request, and then returns to the caller in user mode.

Use of a command dispatcher allows the O/S to provide services that filter access—e.g. a generalised print function will set PEP registers appropriately and ensure QA Chip ink updates occur.

The LEON also allows supervisor mode code to call user mode code in user mode. There are a number of ways that this functionality can be implemented. It is possible to call the user code without a trap, but to return to supervisor mode requires a trap (and associated latency).

3.6 Boot Procedure

3.6.1 Basic Premise

The intention is to load the Manufacturer/owner and OEM program code into SoPEC's RAM, where it can be subsequently executed. The basic SoPEC therefore, must be capable of downloading program code. However SoPEC must be able to guarantee that only authorized Manufacturer/owner boot programs can be loaded, otherwise anyone could modify the O/S to do anything, and then load that—thereby bypassing the licensed operating parameters.

We perform authentication of program code and data using asymmetric (public-key) digital signatures and without using a QA Chip.

Assuming we have already downloaded some data and a 160-bit signature into eDRAM, the boot loader needs to perform the following tasks:

-   -   perform SHA-1 on the downloaded data to calculate a digest         localDigest     -   perform asymmetric decryption on the downloaded signature         (160-bits) using an asymmetric public key to obtain         authorizedDigest     -   If authorizedDigest is the PKCS#1 (patent free) form of         localDigest, then the downloaded data is authorized (the         signature must have been signed with the asymmetric private key)         and control can then be passed to the downloaded data

Asymmetric decryption is used instead of symmetric decryption because the decrypting key must be held in SoPEC's ROM. If symmetric private keys are used, the ROM can be probed and the security is compromised.

The procedure requires the following data item:

-   -   boot0key=an n-bit asymmetric public key

The procedure also requires the following two functions:

-   -   SHA-1=a function that performs SHA-1 on a range of memory and         returns a 160-bit digest     -   decrypt=a function that performs asymmetric decryption of a         message using the passed-in key

PKCS#1 form of localDigest is 2048-bits formatted as follows: bits 2047-2040 = 0x00, bits 2039-2032 = 0x01, bits 2031-288 = 0xFF . . . 0xFF, bits 287-160 = 0x003021300906052B0E03021A05000414, bits 159-0 = localDigest. For more information, see PKCS#1 v2.1 section 9.2

Assuming that all of these are available (e.g. in the boot ROM), boot loader 0 can be defined as in the following pseudocode:

bootloader0(data, sig) localDigest

SHA-1(data) authorizedDigest

decrypt(sig, boot0key) expectedDigest = 0x00|0x01|0xFF..0xFF| 0x003021300906052B0E03021A05000414 |localDigest) // “|” = concat If (authorizedDigest == expectedDigest) jump to program code at data-start address// will never return Else // program code is unauthorized EndIf

The length of the key will depend on the asymmetric algorithm chosen. The key must provide the equivalent protection of the entire QA Chip system—if the Manufacturer/owner O/S program code can be bypassed, then it is equivalent to the QA Chip keys being compromised. In fact it is worse because it would compromise Manufacturer/owner operating parameters, OEM operating parameters, and ink authentication by software downloaded off the net (e.g. from some hacker).

In the case of RSA, a 2048-bit key is required to match the 160-bit symmetric-key security of the QA Chip. In the case of ECDSA, a key length of 132 bits is likely to suffice. RSA is convenient because the patent (U.S. Pat. No. 4,405,829) expired in September 2000.

There is no advantage to storing multiple keys in SoPEC and having the external message choose which key to validate against, because a compromise of any key allows the external user to always select that key.

There is also no particular advantage to having the boot mechanism select the key (e.g. one for USB-based booting and one for external ROM booting) a compromise of the external ROM booting key is enough to compromise all the SoPEC systems.

However, there are advantages in having multiple keys present in the boot ROM and having a wire-bonding option on the pads select which of the keys is to be used. Ideally, the pads would be connected within the package, and the selection is not available via external means once the die has ben packaged. This means we can have different keys for different application areas (e.g. different uses of the chip), and if any particular SoPEC key is compromised, the die could be kept constant and only the bonding changed. Note that in the worst case of all keys being compromised, it may be economically feasible to change the boot0key value in SoPEC's ROM, since this is only a single mask change, and would be easy to verify and characterize.

Therefore the entire security of SoPEC is based on keeping the asymmetric private key paired to boot0key secure. The entire security of SoPEC is also based on keeping the program that signs (i.e. authorizes) datasets using the asymmetric private key paired to boot0key secure.

It may therefore be reasonable to have multiple signatures (and hence multiple signature programs) to reduce the chance of a single point of weakness by a rogue employee. Note that the authentication time increases linearly with the number of signatures, and requires a 2048-bit public key in ROM for each signature.

3.6.2 Hierarchies of Authentication

Given that test programs, evaluation programs, and Manufacturer/owner O/S code needs to be written and tested, and OEM program code etc. also needs to be tested, it is not secure to have a single authentication of a monolithic dataset combining Manufacturer/owner O/S, non-O/S, and OEM program code—we certainly don't want OEMs signing Manufacturer/owner program code, and Manufacturer/owner shouldn't have to be involved with the signing of OEM program code.

Therefore we require differing levels of authentication and therefore a number of keys, although the procedure for authentication is identical to the first—a section of program code contains the key and procedure for authenticating the next.

This method allows for any hierarchy of authentication, based on a root key of boot0key. For example, assume that we have the following entities:

-   -   QACo, Manufacturer/owner's QA/key company. Knows private version         of boot0key, and owner of security concerns.     -   SoPECCo, Manufacturer/owner's SoPEC hardware/software company.         Supplies SoPEC ASICs and SoPEC O/S printing software to a ComCo.     -   ComCo, a company that assembles Print Engines from SoPECs,         Memjet printheads etc, customizing the Print Engine for a given         OEM according to a license     -   OEM, a company that uses a Print Engine to create a printer         product to sell to the end-users. The OEM would supply the motor         control logic, user interface, and casing.

The levels of authentication hierarchy are as follows:

-   -   QACo writes the boot ROM, agenerates dataset1, consisting of a         boot loader program that loads and validates dataset2 and QACo's         asymmetric public boot1key. QACo signs dataset0 with the         asymmetric private boot0key.     -   SoPECCo generates dataset1, consisting of the print engine         security kernel O/S (which incorporates the security-based         features of the print engine functionality) and the ComCo's         asymmetric public key. Upon a special “formal release” request         from SoPECCo, QACo signs dataset0 with QACo's asymmetric private         boot0key key. The print engine program code expects to see an         operating parameter block signed by the ComCo's asymmetric         private key. Note that this is a special “formal release”         request to by SoPECCo; the procedure for development versions of         the program are described in Section 3.6.3.     -   The ComCo generates dataSet3, consisting of dataset1 plus         dataset2, where dataset2 is an operating parameter block for a         given OEM's print engine licence (according to the print engine         license arrangement) signed with the ComCo's asymmetric private         key. The operating parameter block (dataset2) would contain         valid print speed ranges, a PrintEngineLicenseID, and the OEM's         asymmetric public key. The ComCo can generate as many of these         operating parameter blocks for any number of Print Engine         Licenses, but cannot write or sign any supervisor O/S program         code.     -   The OEM would generate dataset5, consisting of dataset3 plus         dataset4, where dataset4 is the OEM program code signed with the         OEM's asymmetric private key. The OEM can produce as many         versions of dataset5 as it likes (e.g. for testing purposes or         for updates to drivers etc) and need not involve         Manufacturer/owner, QACo, or ComCo in any way.

The relationship is shown below in FIG. 325.

When the end-user uses dataset5, SoPEC itself validates dataset1 via the boot0key mechanism described in Section 3.6.1. Once dataset1 is executing, it validates dataset2, and uses dataset2 data to validate dataset4. The validation hierarchy is shown in FIG. 326.

If a key is compromised, it compromises all subsequent authorizations down the hierarchy. In the example from above (and as illustrated in FIG. 326) if the OEM's asymmetric private key is compromised, then O/S program code is not compromised since it is above OEM program code in the authentication hierarchy. However if the ComCo's asymmetric private key is compromised, then the OEM program code is also compromised. A compromise of boot0key compromises everything up to SoPEC itself, and would require a mask ROM change in SoPEC to fix.

It is worthwhile repeating that in any hierarchy the security of the entire hierarchy is based on keeping the asymmetric private key paired to boot0key secure. It is also a requirement that the program that signs (i.e. authorizes) datasets using the asymmetric private key paired to boot0key secure.

3.6.3 Developing Program Code at Manufacturer/Owner

The hierarchical boot procedure described in Section 3.6.1 and Section 3.6.2 gives a hierarchy of protection in a final shipped product.

It is also desirable to use a hierarchy of protection during software development within Manufacturer/owner.

For a program to be downloaded and run on SoPEC during development, it will need to be signed. In addition, we don't want to have to sign each and every Manufacturer/owner development code with the boot0key, as it creates the possibility of any developmental (including buggy or rogue) application being run on any SoPEC.

Therefore QACo needs to generate/create a special intermediate boot loader, signed with boot0key, that performs the exact same tasks as the normal boot loader, except that it checks the SoPECid to see if it is a specific SoPECid (or set of SoPECids). If the SoPEC_id is in the valid set, then the developmental boot loader validates dataset2 by means of its length and a SHA-1 digest of the developmental code³, and not by a further digital signature. The QACo can give this boot loader to the software development team within Manufacturer/owner. The software team can now write and run any program code, and load the program code using the development boot loader. There is no requirement for the subsequent software program (i.e. the developmental program code) to be signed with any key since the programs can only be run on the particular SoPECs. ³The SHA-1 digest is to allow the total program load time to simulate the running time of the normal boot loader running on a non-developmental version of the program.

If the developmental boot loader (and/or signature generator) were compromised, or any of the developmental programs were compromised, the worst situation is that an attacker could run programs on that particular set of SoPECs, and on no others.

This should greatly reduce the possibility of erroneous programs signed with boot0key being available to an attacker (only official releases are signed by boot0key), and therefore reduces the possibility of a Manufacturer/owner employee intentionally or inadvertently creating a back door for attackers.

The relationship is shown below in FIG. 327.

Theoretically the same kind of hierarchy could also be used to allow OEMs to be assured that their program code will only work on specific SoPECs, but this is unlikely to be necessary, and is probably undesirable.

3.6.4 Date-Limited Loaders

It is possible that errors in supervisor program code (e.g. the operating system) could allow attackers to subvert the program in SoPEC and gain supervisor control.

To reduce the impact of this kind of attack, it is possible to allocate some bits of the SoPEC_id to form some kind of date. The granularity of the date could be as simple as a single bit that says the date is obtained from the regular IBM ECID, or it could be 6 bits that give 10 years worth of 3-month units.

The first step of the program loaded by boot loader 0 could check the SoPEC_id date, and run or refuse to run appropriately. The Manufacturer/owner driver or OS could therefore be limited to run on SoPECs that are manufactured up until a particular date.

This means that the OEM would require a new version of the OS for SoPECs after a particular date, but the new driver could be made to work on all previous versions of SoPEC.

The function simply requires a form of date, whose granularity for working can be determined by agreement with the OEM.

For example, suppose that SoPECs are supplied with 3-month granularity in their date components. Manufacturer/owner could ship a version of the OS that works for any SoPEC of the date (i.e. on any chip), or for all SoPECs manufactured during the year etc. The driver issued the next year could work with all SoPECs up until that years etc. In this way the drivers for a chip will be backwards compatible, but will be deliberately not forwards-compatible. It allows the downloading of a new driver with no problems, but it protects against bugs in one years's driver OS from being used against future SoPECs.

Note that the phasing in of a new OS doesn't have to be at the same time as the hardware. For example, the new OS can come in 3 months before the hardware that it supports. However once the new SoPECs are being delivered, the OEM must not ship the older driver with the newer SoPECs, for the old driver will not work on the newer SoPECs. Basically once the OEM has received the new driver, they should use that driver for all SoPEC systems from that point on (old SoPECs will work with the new driver).

This date-limiting feature would most likely be using a field in the ComCo specified operating parameters, so it allows the SoPEC to use date-checking in addition to additional QA Chip related parameter checking (such as the OEM's PrintEngineLicenseId etc).

A variant on this theme is a date-window, where a start-date and end-date are specified (as relating to SoPEC manufacture, not date of use).

3.6.5 Authenticating Operating Parameters

Operating parameters need to be considered in terms of Manufacturer/owner operating parameters and OEM operating parameters. Both sets of operating parameters are stored on the PRINTER_QA chip (physically located inside the printer). This allows the printer to maintain parameters regardless of being moved to different computers, or a loss/replacement of host O/S drivers etc.

On PRINTER_QA, memory vector M₀ contains the upgradable operating parameters, and memory vectors M₁₊ contains any constant (non-upgradable) operating parameters.

Considering only Manufacturer/owner operating parameters for the moment, there are actually two problems:

-   a. setting and storing the Manufacturer/owner operating parameters,     which should be authorized only by Manufacturer/owner -   b. reading the parameters into SoPEC, which is an issue of SoPEC     authenticating the data on the PRINTER_QA chip since we don't trust     PRINTER_QA.

The PRINTER_QA chip therefore contains the following symmetric keys:

-   -   K₀=PrintEngineLicense_key. This key is constant for all SoPECs         supplied for a given print engine license agreement between an         OEM and a Manufacturer/owner ComCo. K₀ has write permissions to         the Manufacturer/owner upgradeable region of M₀ on PRINTER_QA.     -   K₁=SoPEC_id_key. This key is unique for each SoPEC (see Section         3.1), and is known only to the SoPEC and PRINTER_QA. K, does not         have write permissions for anything.

K₀ is used to solve problem (a). It is only used to authenticate the actual upgrades of the operating parameters. Upgrades are performed using the standard upgrade protocol described in [5], with PRINTER_QA acting as the ChipU, and the external upgrader acting as the ChipS.

K₁ is used by SoPEC to solve problem (b). It is used to authenticate reads of data (i.e. the operating parameters) from PRINTER_QA. The procedure follows the standard authenticated read protocol described in [5], with PRINTER_QA acting as ChipR, and the embedded supervisor software on SoPEC acting as ChipT. The authenticated read protocol [5] requires the use of a 160-bit nonce, which is a pseudo-random number. This creates the problem of introducing pseudo-randomness into SoPEC that is not readily determinable by OEM programs, especially given that SoPEC boots into a known state. One possibility is to use the same random number generator as in the QA Chip (a 160-bit maximal-lengthed linear feedback shift register) with the seed taken from the value in the WatchDogTimer register in SoPEC's timer unit when the first page arrives.

Note that the procedure for verifying reads of data from PRINTER_QA does not rely on Manufacturer/owner's key K₀. This means that precisely the same mechanism can be used to read and authenticate the OEM data also stored in PRINTER_QA. Of course this must be done by Manufacturer/owner supervisor code so that SoPEC_id key is not revealed.

If the OEM also requires upgradable parameters, we can add an extra key to PRINTER_QA, where that key is an OEM_key and has write permissions to the OEM part of M₀.

In this way, K₁ never needs to be known by anyone except the SoPEC and PRINTER_QA.

Each printing SoPEC in a multi-SoPEC system need access to a PRINTER_QA chip that contains the appropriate SoPEC_id_key to validate ink useage and operating parameters. This can be accomplished by a separate PRINTER_QA for each SoPEC, or by adding extra keys (multiple SoPEC_id_keys) to a single PRINTER_QA.

However, if ink usage is not being validated (e.g. if print speed were the only Manufacturer/owner upgradable parameter) then not all SoPECs require access to a PRINTER_QA chip that contains the appropriate SoPEC_id_key. Assuming that OEM program code controls the physical motor speed (different motors per OEM), then the PHI within the first (or only) front-page SoPEC can be programmed to accept (or generate) line sync pulses no faster than a particular rate. If line syncs arrived faster than the particular rate, the PHI would simply print at the slower rate. If the motor speed was hacked to be fast, the print image will appear stretched.

3.6.5.1 Floating Operating Parameters and Dongles

As described in Section 2.1.2, Manufacturer/owner operating parameters include such items as print speed, print quality etc. and are tied to a license provided to an OEM. These parameters are under Manufacturer/owner control. The licensed Manufacturer/owner operating parameters are typically stored in the PRINTER_QA as described in Section 3.6.5.

However there are situations when it is desirable to have a floating upgrade to a license, for use on a printer of the user's choice. For example, OEMs may sell a speed-increase license upgrade that can be plugged into the printer of the user's choice. This form of upgrade can be considered a floating upgrade in that it upgrades whichever printer it is currently plugged into. This dongle is referred to as ADDITIONAL_PRINTER_QA. The software checks for the existence of an ADDITIONAL_PRINTER_QA, and if present the operating parameters are chosen from the values stored on both QA chips.

The basic problem of authenticating the additional operating parameters boils down to the problem that we don't trust ADDITIONAL_PRINTER_QA. Therefore we need a system whereby a given SoPEC can perform an authenticated read of the data in ADDITIONAL_PRINTER_QA.

We should not write the SoPEC_id_key to a key in the ADDITIONAL_PRINTER_QA because:

-   -   then it will be tied specifically to that SoPEC, and the primary         intention of the ADDITIONAL_PRINTER_QA is that it be floatable;     -   the ink cartridge would then not work in another printer since         the other printer would not know the old SoPEC_id_key (knowledge         of the old key is required in order to change the old key to a         new one).     -   updating keys is not power-safe (i.e. if at the user's site,         power is removed mid-update, the ADDITIONAL_PRINTER_QA could be         rendered useless)

The proposed solution is to let ADDITIONAL_PRINTER_QA have two keys:

-   -   K₀=FloatingPrintEngineLicense_key. This key has the same         function as the PrintEngineLicense_key in the PRINTER_QA⁴ in         that K₀ has write permissions to the Manufacturer/owner         upgradeable region of M₀ on ADDITIONAL_PRINTER_QA. ⁴This can be         identical to PrintEngineLicense_key in the PRINTER_QA if it is         desireable (unlikely) that upgraders can function on PRINTER_QAs         as well as ADDITIONAL_PRINTER_QAs     -   K₁=UseExtParmsLicense_key. This key is constant for all of the         ADDITIONAL_PRINTER_QAs for a given license agreement between an         OEM and a Manufacturer/owner ComCo (this is not the same key as         PrintEngineLicense_key which is stored as K₀ in PRINTER_QA). K₁         has no write permissions to anything.

K₀ is used to allow writes to the various fields containing operating parameters in the ADDITIONAL_PRINTER_QA. These writes/upgrades are performed using the standard upgrade protocol described in [5], with ADDITIONAL_PRINTER_QA acting as the ChipU, and the external upgrader acting as the ChipS. The upgrader (ChipS) also needs to check the appropriate licensing parameters such as OEM_Id for validity.

K₁ is used to allow SoPEC to authenticate reads of the ink remaining and any other ink data. This is accomplished by having the same UseExtParmsLicense_key within PRINTER_QA (e.g. in K₂), also with no write permissions. i.e:

-   -   PRINTER_QA.K₂=UseExtParmsLicense_key. This key is constant for         all of the PRINTER_QAs for a given license agreement between an         OEM and a Manufacturer/owner ComCo. K₂ has no write permissions         to anything.

This means there are two shared keys, with PRINTER_QA sharing both, and thereby acting as a bridge between INK_CA and SoPEC.

-   -   UseExtParmsLicense_key is shared between PRINTER_QA and         ADDITIONAL_PRINTER_QA     -   SoPEC_id_key is shared between SoPEC and PRINTER_QA

All SoPEC has to do is do an authenticated read [6] from ADDITIONAL_PRINTER_QA, pass the data/signature to PRINTER_QA, let PRINTER_QA validate the data/signature, and get PRINTER_QA to produce a similar signature based on the shared SoPEC_id_key. It can do so using the Translate function [6]. SoPEC can then compare PRINTER_QA's signature with its own calculated signature (i.e. implement a Test function [6] in software on SoPEC), and if the signatures match, the data from ADDITIONAL_PRINTER_QA must be valID, and can therefore be trusted.

Once the data from ADDITIONAL_PRINTER_QA is known to be trusted, the various operating parameters such as OEM_Id can be checked for validity.

The actual steps of read authentication as performed by SoPEC are:

R_(PRINTER)

PRINTER_QA.random( ) R_(DONGLE),M_(DONGLE),SIG_(DONGLE)

DONGLE_QA.read(K1, R_(PRINTER)) R_(SOPEC)

random( ) R_(PRINTER), SIG_(PRINTER)

PRINTER_QA.translate(K2, R_(DONGLE), M_(DONGLE), SIG_(DONGLE), K1, R_(SOPEC)) SIG_(SOPEC)

HMAC_SHA_1(SoPEC_id_key, M_(DONGLE) | R_(PRINTER) | R_(SOPEC)) If (SIG_(PRINTER) = SIG_(SOPEC)) //  various parms inside M_(DONGLE) (data read from ADDITIONAL_PRINTER_QA) is valid Else // the data read from ADDITIONAL_PRINTER_QA is not valid and cannot be trusted EndIf 3.6.5.2 Dongles Tied to a Given SoPEC

Section 3.6.5.1 describes floating dongles i.e. dongles that can be used on any SoPEC. Sometimes it is desirable to tie a dongle to a specific SoPEC.

Tying a QA_CHIP to be used only on a specific SoPEC can be easily accomplished by writing the PRINTER_QA's chipId (unique serial number) into an appropriate M₀ field on the ADDITIONAL_PRINTER_QA. The system software can detect the match and function appropriately. If there is no match, the software can ignore the data read from the ADDITIONAL_PRINTER_QA.

Although it is also possible to store the SoPEC_id_key in one of the keys within the dongle, this must be done in an environment where power will not be removed partway through the key update process (if power is removed during the key update there is a possibility that the dongle QA Chip may be rendered unusable, although this can be checked for after the power failure).

3.6.5.3 OEM Assembly-Line Test

Although an OEM should only be able sell the licensed operating parameters for a given Print Engine, they must be able to assembly-line test⁵ or service/test the Print Engine with a different set of operating parameters e.g. a maximally upgraded Print Engine. ⁵This section is referring to assembly-line testing rather than development testing. An OEM can maximally upgrade a given Print Engine to allow developmental testing of their own OEM program code & mechanics.

Several different mechanisms can be employed to allow OEMs to test the upgraded capabilities of the Print Engine. At present it is unclear exactly what kind of assembly-line tests would be performed.

The simplest solution is to use an ADDITIONAL_PRINTER_QA (i.e. special dongle PRINTER_QA as described in Section 3.6.5.1). The ADDITIONAL_PRINTER_QA would contain the operating parameters that maximally upgrade the printer as long as the dongle is connected to the SoPEC. The exact connection may be directly electrical (e.g. via the standard QA Chip connections) or may be over the USB connection to the printer test host depending on the nature of the test. The exact preferred connection is yet to be determined.

In the testing environment, the ADDITIONAL_PRINTER_QA also requires a numberOfImpressions field inside M₀, which is writeable by K₀. Before the SoPEC prints a page at the higher speed, it decrements the numberOfImpressions counter, performs an authenticated read to ensure the count was decremented, and then prints the page. In this way, the total number of pages that can be printed at high speed is reduced in the event of someone stealing the ADDITIONAL_PRINTER_QA device. It also means that multiple test machines can make use of the same ADDITIONAL_PRINTER_QA.

3.6.6 Use of a PrintEngineLicense id

Manufacturer/owner O/S program code contains the OEM's asymmetric public key to ensure that the subsequent OEM program code is authentic—i.e. from the OEM. However given that SoPEC only contains a single root key, it is theoretically possible for different OEM's applications to be run identically physical Print Engines i.e. printer driver for OEM, run on an identically physical Print Engine from OEM₂.

To guard against this, the Manufacturer/owner O/S program code contains a PrintEngineLicense_id code (e.g. 16 bits) that matches the same named value stored as a fixed operating parameter in the PRINTER_QA (i.e. in M₁₊). As with all other operating parameters, the value of PrintEngineLicense_id is stored in PRINTER_QA (and any ADDITIONAL_PRINTER_QA devices) at the same time as the other various PRINTER_QA customizations are being applied, before being shipped to the OEM site.

In this way, the OEMs can be sure of differentiating themselves through software functionality.

3.6.7 Authentication of Ink

The Manufacturer/owner O/S must perform ink authentication [6] during prints. Ink usage authentication makes use of counters in SoPEC that keep an accurate record of the exact number of dots printed for each ink.

The ink amount remaining in a given cartridge is stored in that cartridge's INK QA chip. Other data stored on the INK_QA chip includes ink color, viscosity, Memjet firing pulse profile information, as well as licensing parameters such as OEM_ID, inkType, InkUsageLicense_ID, etc. This information is typically constant, and is therefore likely to be stored in M₁₊ within INK_QA.

Just as the Print Engine operating parameters are validated by means of PRINTER_QA, a given Print Engine license may only be permitted to function with specifically licensed ink. Therefore the software on SoPEC could contain a valid set of ink types, colors, OEM_Ids, InkUsageLicense_Ids etc. for subsequent matching against the data in the INK_QA.

SoPEC must be able to authenticate reads from the INK_QA, both in terms of ink parameters as well as ink remaining.

To authenticate ink a number of steps must be taken:

-   -   restrict access to dot counts     -   authenticate ink usage and ink parameters via INK_QA and         PRINTER_QA     -   broadcast ink dot usage to all SoPECs in a multi-SoPEC system         3.6.7.1 Restrict Access to Dot Counts

Since the dot counts are accessed via the PHI in the PEP section of SoPEC, access to these registers (and more generally all PEP registers) must be only available from supervisor mode, and not by OEM code (running in user mode). Otherwise it might be possible for OEM program code to clear dot counts before authentication has occurred.

3.6.7.2 Authenticate Ink Usage and Ink Parameters via INK_QA and PRINTER_QA

The basic problem of authentication of ink remaining and other ink data boils down to the problem that we don't trust INK_QA. Therefore how can a SoPEC know the initial value of ink (or the ink parameters), and how can a SoPEC know that after a write to the INK_QA, the count has been correctly decremented.

Taking the first issue, which is determining the initial ink count or the ink parameters, we need a system whereby a given SoPEC can perform an authenticated read of the data in INK_QA.

We cannot write the SoPEC_id_key to the INK_QA for two reasons:

-   -   updating keys is not power-safe (i.e. if power is removed         mid-update, the INK_QA could be rendered useless)     -   the ink cartridge would then not work in another printer since         the other printer would not know the old SoPEC_id_key (knowledge         of the old key is required in order to change the old key to a         new one).

The proposed solution is to let INK_QA have two keys:

-   -   K₀=SupplyInkLicense_key. This key is constant for all ink         cartridges for a given ink supply agreement between an OEM and a         Manufacturer/owner ComCo (this is not the same key as         PrintEngineLicense_key which is stored as K₀ in PRINTER_QA). K₀         has write permissions to the ink remaining regions of M₀ on         INK_QA.     -   K₁=UseInkLicense_key. This key is constant for all ink         cartridges for a given ink usage agreement between an OEM and a         Manufacturer/owner ComCo (this is not the same key as         PrintEngineLicense_key which is stored as K₀ in PRINTER_QA). K₁         has no write permissions to anything.

K₀ is used to authenticate the actual upgrades of the amount of ink remaining (e.g. to fill and refill the amount of ink). Upgrades are performed using the standard upgrade protocol described in [5], with INK_QA acting as the ChipU, and the external upgrader acting as the ChipS. The fill and refill upgrader (ChipS) also needs to check the appropriate ink licensing parameters such as OEM_ID, InkType and InkUsageLicense_Id for validity.

K₁ is used to allow SoPEC to authenticate reads of the ink remaining and any other ink data. This is accomplished by having the same UseInkLicense_key within PRINTER_QA (e.g. in K₂ or K₃), also with no write permissions.

This means there are two shared keys, with PRINTER_QA sharing both, and thereby acting as a bridge between INK_QA and SoPEC.

-   -   UseInkLicense_key is shared between INK_QA and PRINTER_QA     -   SoPEC_id_key is shared between SoPEC and PRINTER_QA

All SoPEC has to do is do an authenticated read [6] from INK_QA, pass the data/signature to PRINTER_QA, let PRINTER_QA validate the data/signature and get PRINTER_QA to produce a similar signature based on the shared SoPEC id key (i.e. the Translate function [6]). SoPEC can then compare PRINTER_QA's signature with its own calculated signature (i.e. implement a Test function [6] in software on the SoPEC), and if the signatures match, the data from INK_QA must be valID, and can therefore be trusted.

Once the data from INK_QA is known to be trusted, the amount of ink remaining can be checked, and the other ink licensing parameters such as OEM_ID, InkType, InkUsageLicense_Id can be checked for validity.

The actual steps of read authentication as performed by SoPEC are:

R_(PRINTER)

PRINTER_QA.random( ) R_(INK), M_(INK), SIG_(INK)

INK_QA.read(K1, R_(PRINTER)) // read with key1: UseInkLicense_key R_(SOPEC)

random( ) R_(PRINTER), SIG_(PRINTER)

PRINTER_QA.translate (K2, R_(INK), M_(INK), SIG_(INK), K1, R_(SOPEC)) SIG_(SOPEC)

HMAC_SHA_1(SoPEC_id_key, M_(INK) | R_(printer) | R_(SOPEC)) If (SIG_(PRINTER) = SIG_(SOPEC)) // M_(INK) (data read from INK_QA) is valid // M_(INK) could be ink parameters, such as InkUsageLicense_Id, or ink remaining If (M_(INK).inkRemaining = expectedInkRemaining) // all is ok Else // the ink value is not what we wrote, so don't print anything anymore EndIf Else // the data read from INK_QA is not valid and cannot be trusted EndIf

Strictly speaking, we don't need a nonce (R_(SOPEC)) all the time because M_(A) (containing the ink remaining) should be decrementing between authentications. However we do need one to retrieve the initial amount of ink and the other ink parameters (at power up). This is why taking a random number from the WatchDogTimer at the receipt of the first page is acceptable.

In summary, the SoPEC performs the non-authenticated write [6] of ink remaining to the INK_QA chip, and then performs an authenticated read of the data via the PRINTER_QA as per the pseudocode above. If the value is authenticated, and the INK_QA ink-remaining value matches the expected value, the count was correctly decremented and the printing can continue.

3.6.7.3 Broadcast Ink Dot Usage to all SoPECs in a Multi-SoPEC System

In a multi-SoPEC system, each SoPEC attached to a printhead must broadcast its ink usage to all the SoPECs. In this way, each SoPEC will have its own version of the expected ink usage.

In the case of a man-in-the-middle attack, at worst the count in a given SoPEC is only its own count (i.e. all broadcasts are turned into 0 ink usage by the man-in-the-middle). We would also require the broadcast amount to be treated as an unsigned integer to prevent negative amounts from being substituted.

A single SoPEC performs the update of ink remaining to the INK_QA chip, and then all SoPECs perform an authenticated read of the data via the appropriate PRINTER_QA (the PRINTER_QA that contains their matching SoPEC_id_key—remember that multiple SoPEC id keys can be stored in a single PRINTER_QA). If the value is authenticated, and the INK_QA value matches the expected value, the count was correctly decremented and the printing can continue.

If any of the broadcasts are not received, or have been tampered with, the updated ink counts will not match. The only case this does not cater for is if each SoPEC is tricked (via a USB2 inter-SoPEC-comms man-in-the-middle attack) into a total that is the same, yet not the true total. Apart from the fact that this is not viable for general pages, at worst this is the maximum amount of ink printed by a single SoPEC. We don't care about protecting against this case.

Since a typical maximum is 4 printing SoPECs, it requires at most 4 authenticated reads. This should be completed within 0.5 seconds, well within the 1-2 seconds/page print time.

3.6.8 Example Hierarchy

Adding an extra bootloader step to the example from Section 3.6.2, we can break up the contents of program space into logical sections, as shown in Table 227. Note that the ComCo does not provide any program code, merely operating parameters that is used by the O/S.

TABLE 227 Sections of Program Space section contents verifies 0 boot loader 0 section 1 via boot0key (ROM) SHA-1 function asymmetric decrypt function boot0key 1 boot loader 1 section 2 via SoPEC_OS_public_key SoPEC_OS_public_key 2 Manufacturer/owner section 3 via O/S program code ComCo_public_key function to generate section 4 via SoPEC_id_key from OEM_public_key (supplied SoPEC_id Basic Print in section 3) Engine PRINTER_QA data, which ComCo_public_key includes the PrintEngineLicense_id, Manufacturer/owner operating parameters, and OEM operating parameters (all authenticated via SoPEC_id_key) 3 ComCo license Is used by section 2 agreement operat- to verify section 4 ing parameter ranges, and range of including parameters as found in PrintEngineLicense_id PRINTER_QA (gets loaded into supervisor mode sec- tion of memory) OEM_public_key (gets loaded into supervisor mode section of memory) Any ComCo written user-mode program code (gets loaded into mode mode section of memory) 4 OEM specific program OEM operating code parameters via calls to Manufacturer/owner O/S code

The verification procedures will be required each time the CPU is woken up, since the RAM is not preserved.

3.6.9 What if the CPU is Not Fast Enough?

In the example of Section 3.6.8, every time the CPU is woken up to print a document it needs to perform:

-   -   SHA-1 on all program code and program data     -   4 sets of asymmetric decryption to load the program code and         data     -   1 HMAC-SHA1 generation per 512-bits of Manufacturer/owner and         OEM printer and ink operating parameters

Although the SHA-1 and HMAC process will be fast enough on the embedded CPU (the program code will be executing from ROM), it may be that the asymmetric decryption will be slow. And this becomes more likely with each extra level of authentication. If this is the case (as is likely), hardware acceleration is required.

A cheap form of hardware acceleration takes advantage of the fact that in most cases the same program is loaded each time, with the first time likely to be at power-up. The hardware acceleration is simply data storage for the authorizedDigest which means that the boot procedure now is:

slowCPU_bootloader0(data, sig) localDigest

SHA-1(data) If (localDigest = previouslyStoredAuthorizedDigest) jump to program code at data-start address// will never return Else authorizedDigest

decrypt(sig, boot0key) expectedDigest = 0x00 |0x01|0xFF..0xFF| Ox003021300906052B0E03021A05000414 | localDigest) If (authorizedDigest == expectedDigest) previouslyStoredAuthorizedDigest

localDigest jump to program code at data-start address// will never return Else // program code is unauthorized EndIf

This procedure means that a reboot of the same authorized program code will only require SHA-1 processing. At power-up, or if new program code is loaded (e.g. an upgrade of a driver over the internet), then the full authorization via asymmetric decryption takes place. This is because the stored digest will not match at power-up and whenever a new program is loaded.

The question is how much preserved space is required.

Each digest requires 160 bits (20 bytes), and this is constant regardless of the asymmetric encryption scheme or the key length. While it is possible to reduce this number of bits, thereby sacrificing security, the cost is small enough to warrant keeping the full digest.

However each level of boot loader requires its own digest to be preserved. This gives a maximum of 20 bytes per loader. Digests for operating parameters and ink levels may also be preserved in the same way, although these authentications should be fast enough not to require cached storage.

Assuming SoPEC provides for 12 digests (to be generous), this is a total of 240 bytes. These 240 bytes could easily be stored as 60×32-bit registers, or probably more conveniently as a small amount of RAM (eg 0.25-1 Kbyte). Providing something like 1 Kbyte of RAM has the advantage of allowing the CPU to store other useful data, although this is not a requirement.

In general, it is useful for the boot ROM to know whether it is being started up due to power-on reset, GPIO activity, or activity on the USB2. In the former case, it can ignore the previously stored values (either 0 for registers or garbage for RAM). In the latter cases, it can use the previously stored values. Even without this, a startup value of 0 (or garbage) means the digest won't match and therefore the authentication will occur implictly.

3.7 SoPEC Phsyical Identification

There must be a mapping of logical to physical since specific SoPECs are responsible for printing on particular physical parts of the page, and/or have particular devices attached to specific pins.

The identification process is mostly solved by general USB2 enumeration.

Each slave SoPEC will need to verify the boot broadcast messages received over USB2, and only execute the code if the signatures are valid. Several levels of authorization may occur. However, at some stage, this common program code (broadcast to all of the slave SoPECs and signed by the appropriate asymmetric private key) can, among other things, set the slave SoPEC's id relating to the physical location. If there is only 1 slave, the id is easy to determine, but if there is more than 1 slave, the id must be determined in some fashion. For example, physical location/id determination may be:

-   -   given by the physical USB2 port on the master     -   related to the physical wiring up of the USB2 interconnects     -   based on GPIO wiring. On other systems, a particular physical         arrangement of SoPECs may exist such that each slave SoPEC will         have a different set of connections on GPIOs. For example, one         SoPEC maybe in charge of motor control, while another may be         driving the LEDs etc. The unused GPIO pins (not necessarily the         same on each SoPEC) can be set as inputs and then tied to 0         or 1. As long as the connection settings are mutually exclusive,         program code can determine which is which, and the id         appropriately set.

This scheme of slave SoPEC identification does not introduce a security breach. If an attacker rewires the pinouts to confuse identification, at best it will simply cause strange printouts (e.g. swapping of printout data) to occur, while at worst the Print Engine will simply not function.

3.8 Setting Up QA Chip Keys

In use, each INK_QA chip needs the following keys:

-   -   K₀=SupplyInkLicense_key     -   K₁=UseInkLicense_key

Each PRINTER_QA chip tied to a specific SoPEC requires the following keys:

-   -   K₀=PrintEngineLicense_key     -   K₁=SoPEC_id_key     -   K₂=UseExtParmsLicense_key     -   K₃=UseInkLicense_key

Note that there may be more than one K₁ depending on the number of PRINTER_QA chips and SoPECs in a system. These keys need to be appropriately set up in the QA Chips before they will function correctly together.

3.8.1 Original QA Chips as Received by a ComCo

When original QA Chips are shipped from QACo to a specific ComCo their keys are as follows:

-   -   K₀=QACo_ComCo_Key0     -   K₁=QACo_ComCo_Key1     -   K₂=QACo_ComCo_Key2     -   K₃=QACo_ComCo_Key3

All 4 keys are only known to QACo. Note that these keys are different for each QA Chip.

3.8.2 Steps at the ComCo

The ComCo is responsible for making Print Engines out of Memjet printheads, QA Chips, PECs or SoPECs, PCBs etc.

In addition, the ComCo must customize the INK_QA chips and PRINTER_QA chip on-board the print engine before shipping to the OEM.

There are two stages:

-   -   replacing the keys in QA Chips with specific keys for the         application (i.e. INK_QA and PRINTER_QA)     -   setting operating parameters as per the license with the OEM         3.8.2.1 Replacing Keys

The ComCo is issued QID hardware [4] by QACo that allows programming of the various keys (except for K₁) in a given QA Chip to the final values, following the standard ChipF/ChipP replace key (indirect version) protocol [6]. The indirect version of the protocol allows each QACo_ComCo_Key to be different for each SoPEC.

In the case of programming of PRINTER_QA's K₁ to be SoPEC_id_key, there is the additional step of transferring an asymmetrically encrypted SoPEC_id_key (by the public-key) along with the nonce (R_(P)) used in the replace key protocol to the device that is functioning as a ChipF. The ChipF must decrypt the SoPEC_id_key so it can generate the standard replace key message for PRINTER_QA (functioning as a ChipP in the ChipF/ChipP protocol). The asymmetric key pair held in the ChipF equivalent should be unique to a ComCo (but still known only by QACo) to prevent damage in the case of a compromise.

Note that the various keys installed in the QA Chips (both INK_QA and PRINTER_QA) are only known to the QACo. The OEM only uses QIDs and QACo supplied ChipFs. The replace key protocol [6] allows the programming to occur without compromising the old or new key.

3.8.2.2 Setting Operating Parameters

There are two sets of operating parameters stored in PRINTER_QA and INK_QA:

-   -   fixed     -   upgradable

The fixed operating parameters can be written to by means of a non-authenticated writes [6] to M₁₊ via a QID [4], and permission bits set such that they are ReadOnly.

The upgradable operating parameters can only be written to after the QA Chips have been programmed with the correct keys as per Section 3.8.2.1. Once they contain the correct keys they can be programmed with appropriate operating parameters by means of a QID and an appropriate ChipS (containing matching keys).

Authentication Protocols

1 Introduction

The following describes authentication protocols for general authentication applications, but with specific reference to the QA Chip.

The intention is to show the broad form of possible protocols for use in different authentication situations, and can be used as a reference when subsequently defining an implementation specification for a particular application. As mentioned earlier, although the protocols are described in relation to a printing environment, many of them have wider application such as, but not limited to, those described at the end of this specification.

2 Nomenclature

The following symbolic nomenclature is used throughout this document:

TABLE 228 Summary of symbolic nomenclature Symbol Description F[X] Function F, taking a single parameter X F[X, Y] Function F, taking two parameters, X and Y X | Y X concatenated with Y X

Y Bitwise X AND Y X

Y Bitwise X OR Y (inclusive-OR) X ⊕ Y Bitwise X XOR Y (exclusive-OR)

X Bitwise NOT X (complement) X

Y X is assigned the value Y X

{Y, Z} The domain of assignment inputs to X is Y and Z X = Y X is equal to Y X ≠ Y X is not equal to Y

X Decrement X by 1 (floor 0)

X Increment X by 1 (modulo register length) Erase X Erase Flash memory register X SetBits[X, Y] Set the bits of the Flash memory register X based on Y Z

ShiftRight[X, Y] Shift register X right one bit position, taking input bit from Y and placing the output bit in Z 3 Pseudocode 3.1 Asynchronous

The following pseudocode:

-   -   var=expression         -   means the var signal or output is equal to the evaluation of             the expression.             3.2 Synchronous

The following pseudocode:

-   -   var←expression         -   means the var register is assigned the result of evaluating             the expression during this cycle.             3.3 Expression

Expressions are defined using the nomenclature in Table 228 above. Therefore:

-   -   var=(a=b)         is interpreted as the var signal is 1 if a is equal to b, and 0         otherwise.         4. Intentionally Blank         5 Basic Protocols         5.1 Protocol Background

This protocol set is a restricted form of a more general case of a multiple key single memory vector protocol. It is a restricted form in that the memory vector M has been optimized for Flash memory utilization:

-   -   M is broken into multiple memory vectors (semi-fixed and         variable components) for the purposes of optimizing flash memory         utilization. Typically M contains some parts that are fixed at         some stage of the manufacturing process (eg a batch number,         serial number etc.), and once set, are not ever updated. This         information does not contain the amount of consumable remaining,         and therefore is not read or written to with any great         frequency.     -   We therefore define M₀ to be the M that contains the frequently         updated sections, and the remaining Ms to be rarely written to.         Authenticated writes only write to M₀, and non-authenticated         writes can be directed to a specific M_(n). This reduces the         size of permissions that are stored in the QA Chip (since         key-based writes are not required for Ms other than M₀). It also         means that M₀ and the remaining Ms can be manipulated in         different ways, thereby increasing flash memory longevity.         5.2 Requirements of Protocol

Each QA Chip contains the following values:

-   N The maximum number of keys known to the chip. -   T The number of vectors M is broken into. -   K_(N) Array of N secret keys used for calculating F_(Kn)[X] where     K_(n) is the nth element of the array. -   R Current random number used to ensure time varying messages. Each     chip instance must be seeded with a different initial value. Changes     for each signature generation. -   M_(T) Array of T memory vectors. Only M₀ can be written to with an     authorized write, while all Ms can be written to in an unauthorized     write. Writes to M₀ are optimized for Flash usage, while updates to     any other M₁₊ are expensive with regards to Flash utilization, and     are expected to be only performed once per section of M_(n). M₁     contains T, N and f in ReadOnly form so users of the chip can know     these two values. -   P_(T+N) T+N element array of access permissions for each part of M.     Entries n={0 . . . T−1} hold access permissions for     non-authenticated writes to M_(n) (no key required). Entries n={T to     T+N−1}hold access permissions for authenticated writes to M₀ for     K_(n). Permission choices for each part of M are Read Only,     Read/Write, and Decrement Only. -   C 3 constants used for generating signatures. C₁, C₂, and C₃ are     constants that pad out a sub-message to a hashing boundary, and all     3 must be different.

Each QA Chip contains the following private function:

-   S_(Kn)[N,X] Internal function only. Returns S_(Kn)[X], the result of     applying a digital signature function S to X based upon the     appropriate key K_(n). The digital signature must be long enough to     counter the chances of someone generating a random signature. The     length depends on the signature scheme chosen, although the scheme     chosen for the QA Chip is HMAC-SHA1, and therefore the length of the     signature is 160 bits.

Additional functions are required in certain QA Chips, but these are described as required.

5.3 Read protocols

The set of read protocols describe the means by which a System reads a specific data vector M_(t) from a QA Chip referred to as ChipR.

We assume that the communications link to ChipR (and therefore ChipR itself) is not trusted. If it were trusted, the System could simply read the data and there is no issue. Since the communications link to ChipR is not trusted and ChipR cannot be trusted, the System needs a way of authenticating the data as actually being from a real ChipR.

Since the read protocol must be capable of being implemented in physical QA Chips, we cannot use asymmetric cryptography (for example the ChipR signs the data with a private key, and System validates the signature using a public key).

This document describes two read protocols:

-   -   direct validation of reads     -   indirect validation of reads.         5.3.1 Direct Validation of Reads

In a direct validation read protocol we require two QA Chips: ChipR is the QA Chip being read, and ChipT is the QA Chip we entrust to tell us whether or not the data read from ChipR is trustworthy. The basic idea is that system asks ChipR for data, and ChipR responds with the data and a signature based on a secret key. System then asks ChipT whether the signature supplied by ChipR is correct. If ChipT responds that it is, then System can trust that data just read from ChipR. Every time data is read from ChipR, the validation procedure must be carried out.

Direct validation requires the System to trust the communication line to ChipT. This could be because ChipT is in physical proximity to the System, and both System and ChipT are in a trusted (e.g. Silverbrook secure) environment. However, since we need to validate the read, ChipR by definition must be in a non-trusted environment.

Each QA Chip protects its signature generation or verification mechanism by the use of a nonce.

The protocol requires the following publicly available functions in ChipT:

-   Random     Returns R (does not advance R). -   Test[n,X, Y, Z] Advances R and returns 1 if S_(Kn)[R|X|C₁|Y]=Z.     Otherwise returns 0. The time taken to calculate and compare     signatures must be independent of data content.

The protocol requires the following publicly available functions in ChipR:

-   Read[n, t, X] Advances R, and returns R, M_(t),     S_(Kn)[X|R|C₁|M_(t)]. The time taken to calculate the signature must     not be based on the contents of X, R, M_(t), or K. If t is invalID,     the function assumes t=0.

To read ChipR's memory M_(t) in a validated way, System performs the following tasks:

-   a. System calls ChipT's Random function; -   b. ChipT returns R_(T) to System; -   c. System calls ChipR's Read function, passing in some key number     n1, the desired data vector number t, and R_(T) (from b); -   d. ChipR updates R_(R), then calculates and returns R_(R), M_(Rt),     S_(Kn1)[R_(T)|R_(R)|C₁|M_(Rt)]; -   e. System calls ChipT's Test function, passing in the key to use for     signature verification n2, and the results from d (i.e. R_(R),     M_(Rt), S_(Kn1)[R_(T)|R_(R)|C₁|M_(Rt)]); -   f. System checks response from ChipT. If the response is 1, then the     M_(t) read from ChipR is considered to be valid. If 0, then the     M_(t) read from ChipR is considered to be invalid.

The choice of n1 and n2 must be such that ChipR's K_(n1)=ChipT's K_(n2).

The data flow for this read protocol is shown in FIG. 328.

From the System's perspective, the protocol would take on a form like the following pseudocode:

R_(T)

ChipT.Random( ) R_(R), M_(R), SIG_(R)

ChipR.Read(keyNumOnChipR,desiredM, R_(T)) ok

ChipT.Test(keyNumOnChipT, R_(R), M_(R), SIG_(R)) If (ok = 1) // M_(R) is to be trusted Else // M_(R) is not to be trusted EndIf

With regards to security, if an attacker finds out ChipR's K_(n1), they can replace the ChipR by a fake ChipR because they can create signatures. Likewise, if an attacker finds out ChipT's K_(n2), they can replace the ChipR by a fake ChipR because ChipR's K_(n1)=ChipT's K_(n2). Moreover, they can use the ChipRs on any system that shares the same key.

The only way of restricting exposure due to key reveals is to restrict the number of systems that match ChipR and ChipT. i.e. vary the key as much as possible. The degree to which this can be done will depend on the application. In the case of a PRINTER_QA acting as a ChipT, and an INK_QA acting as a ChipR, the same key must be used on all systems where the particular INK_QA data must be validated.

In all cases, ChipR must contain sufficient information to produce a signature. Knowing (or finding out) this information, whatever form it is in, allows clone ChipRs to be built.

5.3.2 Indirect Validation of Reads

In a direct validation protocol (see Section 5.3.1), the System validates the correctness of data read from ChipR by means of a trusted chip ChipT. This is possible because ChipR and ChipT share some secret information.

However, it is possible to extend trust via indirect validation. This is required when we trust ChipT, but ChipT doesn't know how to validate data from ChipR. Instead, ChipT knows how to validate data from ChipI (some intermediate chip) which in turn knows how to validate data from either another ChipI (and so on up a chain) or ChipR. Thus we have a chain of validation.

The means of validation chains is translation of signatures. ChipI_(n) translates signatures from higher up the chain (either ChipI_(n−1) or from ChipR at the start of the chain) into signatures capable of being passed to the next stage in the chain (either ChipI_(n+1) or to ChipT at the end of the chain). A given ChipI can only translate signatures if it knows the key of the previous stage in the chain as well as the key of the next stage in the chain.

The protocol requires the following publicly available functions in ChipI:

-   Random     Returns R (does not advance R). -   Translate[n1,X, Y, Z,n2,A] Returns 1, S_(Kn2)[A|R|C₁|Y] and advances     R if Z=S_(Kn1)[R|X|C₁|Y]. Otherwise returns 0, 0. The time taken to     calculate and compare signatures must be independent of data     content.

The data flow for this signature translation protocol is shown in FIG. 329:

Note that R_(prev) is eventually R_(R), and R_(next) is eventually R_(T). In the multiple ChipI case, R_(prev) is the R_(I) of ChipI_(n−1) and R_(next) is R_(I) of ChipI_(n+1). The R_(prev) of the first ChipI in the chain is R_(R), and the R_(next) of the last ChipI in the chain is R_(T).

Assuming at least 1 ChipT, the System would need to perform the following tasks in order to read ChipR's memory M_(t) in an indirectly validated way:

-   a. System calls ChipI_(n)'s Random function; -   b. ChipI₀ returns R₁ to System; -   c. System calls ChipR's Read function, passing in some key number     n0, the desired data vector number t, and R_(I0) (from b); -   d. ChipR updates R_(R), then calculates and returns R_(R), M_(Rt),     S_(Kn0)[R_(In)|R_(R)|C₁|M_(Rt)]; -   e. System assigns R_(R) to R_(prev) and     S_(Kn0)[R_(In)|R_(R)|C₁|M_(Rt)] to SIG_(prev) -   f. System calls the next-chip-in-the-chain's Random function (either     ChipI_(n+), or ChipT) -   g. The next-chip-in-the-chain will return R_(next) to System -   h. System calls ChipI_(n)'s Translate function, passing in n1_(n)     (translation input key number), R_(prev), M_(Rt), SIG_(prev)),     n2_(n), (translation output key number) and the results from g     (R_(next)); -   i. ChipI returns testResult and SIG₁ to System -   j. If testResult=0, then the validation has failed, and the M_(t)     read from ChipR is considered to be invalid. Exit with failure. -   k. If the next chip in the chain is a ChipI, assign SIG₁ to     SIG_(prev) and go to step f -   l. System calls ChipT's Test function, passing in n_(t), R_(prev),     M_(Rt), and SIG_(prev); -   m. System calls System checks response from ChipT. If the response     is 1, then the M_(t) read from ChipR is considered to be valid. If     0, then the M_(t) read from ChipR is considered to be invalid.

For the Translate function to work, ChipI_(n) and ChipI_(n+1), must share a key. The choice of n1 and n2 in the protocol described must be such that ChipI_(n)'s K_(n2)=ChipI_(n+1)'s K_(n1).

Note that Translate is essentially a “Test plus resign” function. From an implementation point of view the first part of Translate is identical to Test.

Note that the use of ChipIs and the translate function merely allows signatures to be transformed. At the end of the translation chain (if present) will be a ChipT requiring the use of a Test function.

There can be any number of ChipIs in the chain to ChipT as long as the Translate function is used to map signatures between ChipI_(n) and ChipI_(n+1) and so on until arrival at the final destination (ChipT).

From the System's perspective, a read protocol using at least 1 ChipI would take on a form like the following pseudocode:

R_(next)

ChipI[0].Random( ) R_(prev), M_(R), SIG_(prev)

ChipR.Read(keyNumOnChipR, desiredM, R_(next)) ok = 1 i = 0 while ((i < iMax) AND ok) For i

0 to iMax If (i = iMax) R_(next)

ChipT.Random( ) Else R_(next)

ChipI[i+1].Random( ) EndIf ok, SIG_(prev) 

 ChipI[i].Translate(iKey[i], R_(prev), M_(R), SIG_(prev), oKey[i], R_(next)) R_(prev) = R_(next) If (ok = 0) // M_(R) is not to be trusted EndIf EndFor ok

ChipT.Test(keyNumOnChipT, R_(prev), M_(R), SIG_(prev)) If (ok = 1) // M_(R) is to be trusted Else // M_(R) is not to be trusted EndIf 5.3.3 Additional Comments on Reads

In the Memjet printing environment, certain implementations will exist where the operating parameters are stored in QA Chips. In this case, the system must read the data from the QA Chip using an appropriate read protocol.

If the connection is trusted (e.g. to a virtual QA Chip in software), a generic Read is sufficient. If the connection is not trusted, it is ideal that the System have a trusted ChipT in the form of software (if possible) or hardware (e.g. a QA Chip on board the same silicon package as the microcontroller and firmware). Whether implemented in software or hardware, the QA Chip should contain an appropriate key that is unique per print engine. Such a key setup would allow reads of print engine parameters and also allow indirect reads of consumables (from a consumable QA Chip).

If the ChipT is physically separate from System (e.g. ChipT is on a board connected to System) System must also occasionally (based on system clock for example) call ChipT's Test function with bad data, expecting a 0 response. This is to reduce the possibility of someone inserting a fake ChipT into the system that always returns 1 for the Test function.

5.4 Upgrade Protocols

This set of protocols describe the means by which a System upgrades a specific data vector M_(t) within a QA Chip (ChipU). The data vector may contain information about the functioning of the device (e.g. the current maximum operating speed) or the amount of a consumable remaining.

The updating of M_(t) in ChipU falls into two categories:

-   -   non-authenticated writes, where anyone is able to update the         data vector     -   authenticated writes, where only authorized entities are able to         upgrades data vectors         5.4.1 Non-Authenticated Writes

This is the most frequent type of write, and takes place between the System/consumable during normal everyday operation for M₀, and during the manufacturing process for M₁₊.

In this kind of write, the System wants to change M_(t) within ChipU subject to P. For example, the System could be decrementing the amount of consumable remaining. Although System does not need to know and of the K_(S) or even have access to a trusted chip to perform the write, the System must follow a non-authenticated write by an authenticated read if it needs to know that the write was successful.

The protocol requires ChipU to contain the following publicly available function:

-   Write[t, X] Writes X over those parts of M_(t) subject to P_(t) and     the existing value for M.     -   To authenticate a write of M_(new) to ChipA's memory M_(new): -   a. System calls ChipU's Write function, passing in M_(new); -   b. The authentication procedure for a Read is carried out (see     Section 5.3 on page 604); -   c. If the read succeeds in such a way that M_(new)=M returned in b,     the write succeeded. If not, it failed.

Note that if these parameters are transmitted over an error-prone communications line (as opposed to internally or using an additional error-free transport layer), then an additional checksum would be required to prevent the wrong M from being updated or to prevent the correct M from being updated to the wrong value. For example, SHA-1 [t,X] should be additionally transferred across the communications line and checked (either by a wrapper function around Write or in a variant of Write that takes a hash as an extra parameter).

This is the most frequent type of write, and takes place between the System/consumable during normal everyday operation for M₀, and during the manufacturing process for M₁₊.

5.4.2 Authenticated Writes

In the QA Chip protocols, M₀ is defined to be the only data vector that can be upgraded in an authenticated way. This decision was made primarily to simplify flash management, although it also helps to reduce the permissions storage requirements.

In this kind of write, System wants to change Chip U's M₀ in an authorized way, without being subject to the permissions that apply during normal operation. For example, a consumable may be at a refilling station and the normally Decrement Only section of M₀ should be updated to include the new valid consumable. In this case, the chip whose M₀ is being updated must authenticate the writes being generated by the external System and in addition, apply the appropriate permission for the key to ensure that only the correct parts of M₀ are updated. Having a different permission for each key is required as when multiple keys are involved, all keys should not necessarily be given open access to M₀. For example, suppose M₀ contains printer speed and a counter of money available for franking. A ChipS that updates printer speed should not be capable of updating the amount of money. Since P_(0 . . . T−1) is used for non-authenticated writes, each K_(n) has a corresponding permission P_(T+n) that determines what can be updated in an authenticated write.

The basic principle of the authenticated write (or upgrade) protocol is that the new value for the M_(t) must be signed before ChipU accepts it. The QA Chip responsible for generating the signature (ChipS) must first validate that the ChipU is valid by reading the old value for M_(t). Once the old value is seen as valID, a new value can be signed by ChipS and the resultant data plus signature passed to ChipU. Note that both chips distrust each other.

There are two forms of authenticated writes. The first form is when both ChipU and ChipS directly store the same key. The second is when both ChipU and ChipS store different versions of the key and a transforming procedure is used on the stored key to generate the required key—i.e. the key is indirectly stored. The second form is slightly more complicated, and only has value when the ChipS is not readily available to an attacker.

5.4.2.1 Direct Authenticated Writes

The direct form of the authenticated write protocol is used when the ChipS and ChipU are equally available to an attacker. For example, suppose that ChipU contains a printer's operating speed. Suppose that the speed can be increased by purchasing a ChipS and inserting it into the printer system. In this case, the ChipS and ChipU are equally available to an attacker. This is different from upgrading the printer over the internet where the effective ChipS is in a remote location, and thereby not as readily available to an attacker.

The direct authenticated write protocol requires ChipU to contain the following publicly available functions:

-   Read[n, t, X] Advances R, and returns R, M_(t),     S_(Kn)[X|R|C₁|M_(t)]. The time taken to calculate the signature must     not be based on the contents of X, R, M_(t), or K. -   WriteA[n, X, Y, Z] Advances R, replaces M₀ by Y subject to P_(T+n),     and returns 1 only if S_(Kn)[R|X|C₁|Y]=Z. Otherwise returns 0. The     time taken to calculate and compare signatures must be independent     of data content. This function is identical to ChipT's Test function     except that it additionally writes Y subject to P_(T+n) to its M     when the signature matches.

Authenticated writes require that the System has access to a ChipS that is capable of generating appropriate signatures.

In its basic form, ChipS requires the following variables and function:

-   SignM[n,V,W,X,Y,Z] Advances R, and returns R, S_(Kn)[W|R|C₁|Z] only     if Y=S_(Kn)[V|W|C₁|X]. Otherwise returns all 0s. The time taken to     calculate and compare signatures must be independent of data     content.

To update ChipU's M vector:

-   a. System calls ChipU's Read function, passing in n1, 0 (desired     vector number) and 0 (the random value, but is a don't-care value)     as the input parameters; -   b. ChipU produces R_(U), M_(U0), S_(Kn1)[0|R_(U)|C₁|M_(U0)] and     returns these to System; -   c. System calls ChipS's SignM function, passing in n2 (the key to be     used in ChipS), 0 (the random value as used in a), R_(U), M_(U0),     S_(Kn1)[0|R_(U)|C₁|M_(U0)], and M_(D) (the desired vector to be     written to ChipU); -   d. ChipS produces R_(S) and S_(Kn2)[R_(U)|R_(S)|C₁|M_(D)] if the     inputs were valID, and 0 for all outputs if the inputs were not     valid. -   e. If values returned in d are non zero, then ChipU is considered     authentic. System can then call ChipU's WriteA function with these     values from d. -   f. ChipU should return a 1 to indicate success. A 0 should only be     returned if the data generated by ChipS is incorrect (e.g. a     transmission error).

The choice of n1 and n2 must be such that ChipU's K_(n1)=ChipS's K_(n2).

The data flow for authenticated writes is shown in FIG. 330.

Note that this protocol allows ChipS to generate a signature for any desired memory vector M_(D), and therefore a stolen ChipS has the ability to effectively render the particular keys for those parts of M₀ in ChipU irrelevant.

It is therefore not recommended that the basic form of ChipS be ever implemented except in specifically controlled circumstances.

It is much more secure to limit the powers of ChipS. The following list covers some of the variants of limiting the power of ChipS:

-   a. the ability to upgrade a limited number of times -   b. the ability to upgrade based on a credit value—i.e. the upgrade     amount is decremented from the local value, and effectively     transferred to the upgraded device -   c. the ability to upgrade to a fixed value or from a limited list -   d. the ability to upgrade to any value -   e. the ability to only upgrade certain data fields within M

In many of these variants, the ability to refresh the ChipS in some way (e.g. with a new count or credit value) would be a useful feature.

In certain cases, the variant is in ChipS, while ChipU remains the same. It may also be desirable to create a ChipU variant, for example only allowing ChipU to only be upgraded a specific number of times.

5.4.2.1.1 Variant Example

This section details the variant for the ability to upgrade a memory vector to any value a specific number of times, but the upgrade is only allowed to affect certain fields within the memory vector i.e. a combination of (a), (d), and (e) above.

In this example, ChipS requires the following variables and function:

-   CountRemaining Part of ChipS's M₀ that contains the number of     signatures that ChipS is allowed to generate. Decrements with each     successful call to SignM and SignP. Permissions in ChipS's     P_(0..T−1) for this part of M₀ needs to be ReadOnly once ChipS has     been setup. Therefore CountRemaining can only be updated by another     ChipS that will perform updates to that part of M₀ (assuming ChipS's     Ps allows that part of M₀ to be updated). -   Q Part of M that contains the write permissions for updating     ChipU's M. By adding Q to ChipS we allow different ChipSs that can     update different parts of M_(U). Permissions in ChipS's P_(0..T−1)     for this part of M needs to be ReadOnly once ChipS has been setup.     Therefore Q can only be updated by another ChipS that will perform     updates to that part of M. -   SignM[n,V,W,X,Y,Z] Advances R, decrements CountRemaining and returns     R, Z_(QX) (Z applied to X with permissions Q), S_(Kn)[W|R|C₁|Z_(QX)]     only if Y=S_(Kn)[V|W|C₁|X] and CountRemaining>0. Otherwise returns     all 0s. The time taken to calculate and compare signatures must be     independent of data content.

To update ChipU's M vector:

-   a. System calls ChipU's Read function, passing in n1, 0 (desired     vector number) and 0 (the random value, but is a don't-care value)     as the input parameters; -   b. ChipU produces R_(U), M_(U0), S_(Kn1)[0|R_(U)|C₁|M_(U0)] and     returns these to System; -   c. System calls ChipS's SignM function, passing in n2 (the key to be     used in ChipS), 0 (as used in a), R_(U), M_(U0),     S_(Kn1)[0|R_(U)|C₁|M_(U0)], and M_(D) (the desired vector to be     written to ChipU); -   d. ChipS produces R_(S), M_(QD) (processed by running M_(D) against     M_(U0) using Q) and S_(Kn2)[R_(U)|R_(S)|C₁|M_(QD)] if the inputs     were valID, and 0 for all outputs if the inputs were not valid. -   e. If values returned in d are non zero, then ChipU is considered     authentic. System can then call ChipU's WriteA function with these     values from d. -   f. ChipU should return a 1 to indicate success. A 0 should only be     returned if the data generated by ChipS is incorrect (e.g. a     transmission error).

The choice of n1 and n2 must be such that ChipU's K_(n1)=ChipS's K_(n2).

The data flow for this variant of authenticated writes is shown in FIG. 331.

Note that Q in ChipS is part of ChipS's M. This allows a user to set up ChipS with a permission set for upgrades. This should be done to ChipS and that part of M designated by P_(0..T−1) set to ReadOnly before ChipS is programmed with K_(U). If K_(S) is programmed with K_(U) first, there is a risk of someone obtaining a half-setup ChipS and changing all of M_(U) instead of only the sections specified by Q.

In addition, CountRemaining in ChipS needs to be setup (including making it ReadOnly in P_(S)) before ChipS is programmed with K_(U). ChipS should therefore be programmed to only perform a limited number of SignM operations (thereby limiting compromise exposure if a ChipS is stolen). Thus ChipS would itself need to be upgraded with a new CountRemaining every so often.

5.4.2.2 Indirect Authenticated Writes

This section describes an alternative authenticated write protocol when ChipU is more readily available to an attacker and ChipS is less available to an attacker. We can store different keys on ChipU and ChipS, and implement a mapping between them in such a way that if the attacker is able to obtain a key from a given ChipU, they cannot upgrade all ChipUs.

In the general case, this is accomplished by storing key K_(S) on ChipS, and K_(U) and f on ChipU. The relationship is f(K_(S))=K_(U) such that knowledge of K_(U) and f does not make it easy to determine K_(S). This implies that a one-way function is desirable for f.

In the QA Chip domain, we define f as a number (e.g. 32-bits) such that SHA1(K_(S)|f)=K_(U). The value of f (random between chips) can be stored in a known location within M₁ as a constant for the life of the QA Chip. It is possible to use the same f for multiple relationships if desired, since f is public and the protection lies in the fact that f varies between QA Chips (preferably in a non-predictable way).

The indirect protocol is the same as the direct protocol with the exception that f is additionally passed in to the SignM function so that ChipS is able to generate the correct key. The System obtains f by performing a Read of M₁. Note that all other functions, including the WriteA function in ChipU, are identical to their direct authentication counterparts.

-   SignM[f,n,V,W,X,Y,Z] Advances R, and returns R, S_(f(Kn))[W|R|C₁|Z]     only if Y=S_(f(Kn))[V|W|C₁|X] and CountRemaining>0. Otherwise     returns all 0s. The time taken to calculate and compare signatures     must be independent of data content.

Before reading ChipU's memory M₀ (the pre-upgrade value), the System must extract f from ChipU by performing the following tasks:

-   a. System calls ChipU's Read function, passing in (dontCare, 1,     dontCare) -   b. ChipU returns M₁, from which System can extract f_(U) -   c. System stores f_(U) for future use

To update ChipU's M vector, the protocol is identical to that described in the basic authenticated write protocol with the exception of steps c and d:

-   c. System calls ChipS's SignM function, passing in f_(U), n2 (the     key to be used in ChipS), 0 (as used in a), R_(U), M_(U0),     S_(Kn1)[0|R_(U)|C₁|M_(U0)], and M_(D) (the desired vector to be     written to ChipU); -   d. ChipS produces R_(S) and S_(fU(Kn2))[R_(U)|R_(S)|C₁|M_(D)] if the     inputs were valID, and 0 for all outputs if the inputs were not     valid.

In addition, the choice of n1 and n2 must be such that ChipU's K_(n1)=ChipS's f_(U)(K_(n2)).

Note that f_(U) is obtained from M₁ without validation. This is because there is nothing to be gained by subverting the value of f_(U), (because then the signatures won't match).

From the System's perspective, the protocol would take on a form like the following pseudocode:

dontCare, M_(R), dontCare

ChipR.Read(dontCare,1, dontCare) f_(R) = extract from M_(R) ... R_(U), M_(U), SIG_(U)

ChipU.Read(keyNumOnChipU,0, 0) R_(S), SIG_(S) = ChipS.SignM2(f_(R), keyNumOnChipS, 0, R_(U), M_(U), SIG_(U), M_(D)) If (R_(S) = SIG_(S) = 0) // ChipU and therefore M_(U) is not to be trusted Else // ChipU and therefore M_(U) can be trusted ok = ChipU.WriteA(keyNumOnChipU, R_(S), M_(D), SIG_(S)) If (ok) // updating of data in ChipU was successful Else // transmission error during WriteA EndIf EndIf 5.4.2.2.1 Variant Example

The indirect form of the example from Section 5.4.2.1.1 is shown here.

-   SignM[f,n,V,W,X,Y,Z] Advances R, decrements CountRemaining and     returns R, Z_(QX) (Z applied to X with permissions Q),     S_(f(Kn))[W|R|C₁|Z_(QX)] only if Y=S_(f(Kn))[V|W|C₁X] and     CountRemaining>0. Otherwise returns all 0s. The time taken to     calculate and compare signatures must be independent of data     content.

Before reading ChipU's memory M₀ (the pre-upgrade value), the System must extract f from ChipU by performing the following tasks:

-   a. System calls ChipU's Read function, passing in (dontCare, 1,     dontCare) -   b. ChipU returns M₁, from which System can extract f_(U) -   c. System stores f_(U) for future use

To update ChipU's M vector, the protocol is identical to that described in the basic authenticated write protocol with the exception of steps c and d:

-   c. System calls ChipS's SignM function, passing in f_(U), n2 (the     key to be used in ChipS), 0 (as used in a), R_(U), M_(U0),     S_(Kn1)[0|R_(U)|C₁|M_(U0)], and M_(D) (the desired vector to be     written to ChipU); -   d. ChipS produces R_(S), M_(QD) (processed by running M_(D) against     M_(U0) using Q) and S_(fU(Kn2))[R_(U)|R_(S)|C₁|M_(QD)] if the inputs     were valID, and 0 for all outputs if the inputs were not valid.

In addition, the choice of n1 and n2 must be such that ChipU's K_(n1)=ChipS's f_(U)(K_(n2)).

Note that f_(U) is obtained from M₁ without validation. This is because there is nothing to be gained by subverting the value of f_(U), (because then the signatures won't match).

From the System's perspective, the protocol would take on a form like the following pseudocode:

dontCare, M_(R), dontCare

ChipR.Read(dontCare,1, dontCare) f_(R) = extract from M_(R) ... R_(U), M_(U), SIG_(U)

ChipU.Read(keyNumOnChipU,0, 0) R_(S), M_(QD), SIG_(S) = ChipS.SignM2(f_(R), keyNumOnChipS, 0, R_(U), M_(U), SIG_(U), M_(D)) If (R_(S) = M_(QD) = SIG_(S) = 0) // ChipU and therefore M_(U) is not to be trusted Else // ChipU and therefore M_(U) can be trusted ok = ChipU.WriteA(keyNumOnChipU, R_(S), M_(QD), SIG_(S)) If (ok) // updating of data in ChipU was successful Else // transmission error during WriteA EndIf EndIf 5.4.3 Updating Permissions for Future Writes

In order to reduce exposure to accidental and malicious attacks on P (and certain parts of M), only authorized users are allowed to update P. Writes to P are the same as authorized writes to M, except that they update P_(n) instead of M. Initially (at manufacture), P is set to be Read/Write for all M. As different processes fill up different parts of M, they can be sealed against future change by updating the permissions. Updating a chip's P_(0..T−1) changes permissions for unauthorized writes to M_(n), and updating P_(T..T+N−1) changes permissions for authorized writes with key K_(n).

P_(n) is only allowed to change to be a more restrictive form of itself. For example, initially all parts of M have permissions of Read/Write. A permission of Read/Write can be updated to Decrement Only or Read Only. A permission of Decrement Only can be updated to become Read Only. A Read Only permission cannot be further restricted.

In this transaction protocol, the System's chip is referred to as ChipS, and the chip being updated is referred to as ChipU. Each chip distrusts the other.

The protocol requires the following publicly available functions in ChipU:

-   Random     Returns R (does not advance R). -   SetPermission[n,p,X,Y,Z] Advances R, and updates P_(p) according to     Y and returns 1 followed by the resultant P_(p) only if     S_(Kn)[R|X|Y|C₂]=Z. Otherwise returns 0. P_(p) can only become more     restricted. Passing in 0 for any permission leaves it unchanged     (passing in Y=0 returns the current P_(p)).

Authenticated writes of permissions require that the System has access to a ChipS that is capable of generating appropriate signatures. ChipS requires the following variable:

-   CountRemaining Part of ChipS's M₀ that contains the number of     signatures that ChipS is allowed to generate. Decrements with each     successful call to SignM and SignP. Permissions in ChipS's     P_(0..T−1) for this part of M₀ needs to be ReadOnly once ChipS has     been setup. Therefore CountRemaining can only be updated by another     ChipS that will perform updates to that part of M₀ (assuming ChipS's     P_(n) allows that part of M₀ to be updated).

In addition, ChipS requires either of the following two SignP functions depending on whether direct or indirect key storage is used (see direct vs indirect authenticated write protocols in Section 5.4.2):

-   SignP[n,X,Y] Used when the same key is directly stored in both ChipS     and ChipU. Advances R, decrements CountRemaining and returns R and     S_(Kn)[X|R|Y|C₂] only if CountRemaining>0. Otherwise returns all 0s.     The time taken to calculate and compare signatures must be     independent of data content. -   SignP[f,n,X,Y] Used when the same key is not directly stored in both     ChipS and ChipU. In this case ChipU's K_(n1)=ChipS's f(K_(n2)). The     function is identical to the direct form of SignP, except that it     additionally accepts f and returns S_(f(Kn))[X|R|Y|C₂] instead of     S_(Kn)[X|R|Y|C₂].     5.4.3.1 Direct Form of SignP

When the direct form of SignP is used, ChipU's P_(n) is updated as follows:

-   a. System calls ChipU's Random function; -   b. ChipU returns R_(U) to System; -   c. System calls ChipS's SignP function, passing in n2, R_(U) and     P_(D) (the desired P to be written to ChipU); -   d. ChipS produces R_(S) and S_(Kn2)[R_(U)|R_(S)|P_(D)|C₂] if it is     still permitted to produce signatures. -   e. If values returned in d are non zero, then System can then call     ChipU's SetPermission function with n1, the desired permission entry     p, R_(S), P_(D) and S_(Kn2)[R_(U)|R_(S)|P_(D)|C₂]. -   f. ChipU verifies the received signature against its own generated     signature S_(Kn1)[R_(U)|R_(S)|P_(D)|C₂] and applies P_(D) to P_(n)     if the signature matches -   g. System checks 1 st output parameter. 1=success, 0=failure.

The choice of n1 and n2 must be such that ChipU's K_(n1)=ChipS's K_(n2).

The data flow for basic authenticated writes to permissions is shown in FIG. 332.

5.4.3.2 Indirect Form of SignP

When the indirect form of SignP is used in ChipS, the System must extract f from ChipU (so it knows how to generate the correct key) by performing the following tasks:

-   a. System calls ChipU's Read function, passing in (dontCare, 1,     dontCare) -   b. ChipU returns M₁, from which System can extract f_(U) -   c. System stores f_(U) for future use

ChipU's P_(n) is updated as follows:

-   a. System calls ChipU's Random function; -   b. ChipU returns R_(U) to System; -   c. System calls ChipS's SignP function, passing in f_(U), n2, R_(U)     and P_(D) (the desired P to be written to ChipU); -   d. ChipS produces R_(S) and S_(fU(K) _(n2))[R_(U)|R_(S)|P_(D)|C₂] if     it is still permitted to produce signatures. -   e. If values returned in d are non zero, then System can then call     ChipU's SetPermission function with n1, the desired permission entry     p, R_(S), P_(D) and S_(fU(Kn2))[R_(U)|R_(S)|P_(D)|C₂]. -   f. ChipU verifies the received signature against     S_(Kn1)[R_(U)|R_(S)|P_(D)|C₂] and applies P_(D) to P_(n) if the     signature matches -   g. System checks 1 st output parameter. 1=success, 0=failure.

In addition, the choice of n1 and n2 must be such that ChipU's K_(n1)=ChipS's f_(U)(K_(n2)).

5.4.4 Protecting Memory Vectors

To protect the appropriate part of M_(n) against unauthorized writes, call SetPermissions[n] for n=0 to T−1. To protect the appropriate part of M₀ against authorized writes with key n, call SetPermissions[T+n] for n=0 to N−1.

Note that only M₀ can be written in an authenticated fashion.

Note that the SetPermission function must be called after the part of M has been set to the desired value.

For example, if adding a serial number to an area of M₁ that is currently ReadWrite so that noone is permitted to update the number again:

-   -   the Write function is called to write the serial number to M₁     -   SetPermission(1) is called for to set that part of M to be         ReadOnly for non-authorized writes.

If adding a consumable value to M₀ such that only keys 1-2 can update it, and keys 0, and 3-N cannot:

-   -   the Write function is called to write the amount of consumable         to M     -   SetPermission is called for 0 to set that part of M₀ to be         DecrementOnly for non-authorized writes. This allows the amount         of consumable to decrement.     -   SetPermission is called for n={T, T+3, T+4 . . . , T+N−1} to set         that part of M₀ to be ReadOnly for authorized writes using all         but keys 1 and 2. This leaves keys 1 and 2 with ReadWrite         permissions to M₀.

It is possible for someone who knows a key to further restrict other keys, but it is not in anyone's interest to do so.

5.5 Programming K

In this case, we have a factory chip (ChipF) connected to a System. The System wants to program the key in another chip (ChipP). System wants to avoid passing the new key to ChipP in the clear, and also wants to avoid the possibility of the key-upgrade message being replayed on another ChipP (even if the user doesn't know the key).

The protocol assumes that ChipF and ChipP already share (directly or indirectly) a secret key K_(old). This key is used to ensure that only a chip that knows K_(old) can set K_(new).

Although the example shows a ChipF that is only allowed to program a specific number of ChipPs, the key-upgrade protocol can be easily altered (similar to the way the write protocols have variants) to provide other means of limiting the ability to update ChipPs.

The protocol requires the following publicly available functions in ChipP:

-   Random     Returns R (does not advance R).     -   ReplaceKey[n, X, Y, Z] Replaces K_(n) by S_(Kn)[R|X|C₃]⊕Y,         advances R, and returns 1 only if S_(Kn)[X|Y|C₃]=Z. Otherwise         returns 0. The time taken to calculate signatures and compare         values must be identical for all inputs.

And the following data and functions in ChipF:

-   CountRemaining Part of M₀ with contains the number of signatures     that ChipF is allowed to generate. Decrements with each successful     call to GetProgramKey. Permissions in P for this part of M₀ needs to     be ReadOnly once ChipF has been setup. Therefore can only be updated     by a ChipS that has authority to perform updates to that part of M₀. -   K_(new) The new key to be transferred from ChipF to ChipP. Must not     be visible. After manufacture, K_(new) is 0.

SetPartialKey[X] Updates K_(new) to be K_(new)⊕X. This function allows K_(new) to be programmed in any number of steps, thereby allowing different people or systems to know different parts of the key (but not the whole K_(new)). K_(new) is stored in ChipF's flash memory.

In addition, ChipF requires either of the following GetProgramKey functions depending on whether direct or indirect key storage is used on the input key and/or output key (see direct vs indirect authenticated write protocols in Section 5.4.2):

-   GetProgramKey1[n, X] Direct to direct. Used when the same key     (K_(n)) is directly stored in both ChipF and ChipP and we want to     store K_(new) in ChipP. Advances R_(F), decrements CountRemaining,     outputs R_(F), the encrypted key S_(Kn)[X|R_(F)|C₃]⊕K_(new) and a     signature of the first two outputs plus C₃ if CountRemaining>0.     Otherwise outputs 0. The time to calculate the encrypted key &     signature must be identical for all inputs. -   GetProgramKey2[f, n, X] Direct to indirect. Used when the same key     (K_(n)) is directly stored in both ChipF and ChipP but we want to     store f_(P)(K_(new)) in ChipP instead of simply K_(new) (i.e. we     want to keep the key in ChipP to be different in all ChipPs). In     this case ChipP's K_(n1)=ChipF's f_(P)(K_(n2)). The function is     identical to GetProgramKey1, except that it additionally accepts     f_(P), and returns S_(Kn)[X|R_(F)|C₃]⊕f_(P)(K_(new)) instead of     S_(Kn)[X|R_(F)|C₃]⊕K_(new). Note that the produced signature is     produced using K_(n) since that is what is already stored in ChipP. -   GetProgramKey3[f, n, X] Indirect to direct. Used when the same key     is not directly stored in both ChipF and ChipP but we want to store     K_(new) in ChipP. In this case ChipP's K_(n1)=ChipF's f_(P)(K_(n2)).     The function is identical to GetProgramKey1, except that it     additionally accepts f_(P), and returns     S_(fP(Kn))[X|R_(F)|C₃]⊕K_(new) instead of     S_(Kn)[X|R_(F)|C₃]⊕K_(new). The produced signature is produced using     f_(P)(Kn) instead of K_(n) since that is what is already stored in     ChipP. -   GetProgramKey4[f, n, X] Indirect to indirect. Used when the same key     is not directly stored in both ChipF and ChipP but we want to store     f_(P)(K_(new)) in ChipP instead of simply K_(new) (i.e. we want to     keep the key in ChipP to be different in all ChipPs). In this case     ChipP's K_(n1)=ChipF's f_(P)(K_(n2)). The function is identical to     GetProgramKey3, except that it returns     S_(fP(Kn))[X|R_(F)|C₃]⊕f_(P)(K_(new)) instead of     S_(fP(Kn))[X|R_(F)|C₃]⊕K_(new). The produced signature is produced     using f_(P)(K_(n)) since that is what is already stored in ChipP.

Since there are likely to be few ChipFs, and many ChipPs, the indirect forms of GetProgramKey can be usefully employed.

5.5.1 GetProgramKey1—direct to direct

With the “old key=direct, new key=direct” form of GetProgramKey, to update P's key:

-   a. System calls ChipP's Random function; -   b. ChipP returns R_(P) to System; -   c. System calls ChipF's GetProgramKey function, passing in n2 (the     desired key to use) and the result from b; -   d. ChipF updates R_(F), then calculates and returns R_(F),     S_(Kn2)[R_(P)|R_(F)|C₃]⊕K_(new), and     S_(Kn2)[R_(F)|S_(Kn2)[R_(P)|R_(F)|C₃|⊕K_(new)⊕C₃]; -   e. If the response from d is not 0, System calls ChipP's ReplaceKey     function, passing in n1 (the key to use in ChipP) and the response     from d; -   f. System checks response from ChipP. If the response is 1, then     ChipP's K_(n1) has been correctly updated to K_(new). If the     response is 0, ChipP's K_(n1) has not been updated.

The choice of n1 and n2 must be such that ChipP's K_(n1)=ChipF's K_(n2).

The data flow for key updates is shown in FIG. 333:

Note that K_(new) is never passed in the open. An attacker could send its own R_(P), but cannot produce S_(Kn2)[R_(P)|R_(F)|C₃] without K_(n2). The signature based on K_(new) is sent to ensure that ChipP will be able to determine if either of the first two parameters have been changed en route.

CountRemaining needs to be setup in M_(F0) (including making it ReadOnly in P) before ChipF is programmed with K_(P). ChipF should therefore be programmed to only perform a limited number of GetProgramKey operations (thereby limiting compromise exposure if a ChipF is stolen). An authorized ChipS can be used to update this counter if necessary (see Section 5.4.2 on page 610).

5.5.2 GetProgramKey2—Direct to Indirect

With the “old key=direct, new key=indirect” form of GetProgramKey, to update P's key, the System must extract f from ChipP (so it can tell ChipF how to generate the correct key) by performing the following tasks:

-   a. System calls ChipP's Read function, passing in (dontCare, 1,     dontCare) -   b. ChipP returns M₁, from which System can extract f_(P) -   c. System stores f_(P) for future use

ChipP's key is updated as follows:

-   a. System calls ChipP's Random function; -   b. ChipP returns R_(P) to System; -   c. System calls ChipF's GetProgramKey function, passing in f_(P), n2     (the desired key to use) and the result from b; -   d. ChipF updates R_(F), then calculates and returns R_(F),     S_(Kn2)[R_(P)|R_(F)|C₃]⊕f_(P)(K_(new)), and     S_(Kn2)[R_(F)|S_(Kn2)[R_(P)|R_(F)|C₃]⊕f_(P)(K_(new))|C₃]; -   e. If the response from d is not 0, System calls ChipP's ReplaceKey     function, passing in n1 (the key to use in ChipP) and the response     from d; -   f. System checks response from ChipP. If the response is 1, then     ChipP's K_(n1) has been correctly updated to f_(P)(K_(new)). If the     response is 0, ChipP's K_(n1) has not been updated.

The choice of n1 and n2 must be such that ChipP's K_(n1)=ChipF's K_(n2).

5.5.3 GetProgramKey3—Indirect to Direct

With the “old key=indirect, new key=direct” form of GetProgramKey, to update P's key, the System must extract f from ChipP (so it can tell ChipF how to generate the correct key) by performing the following tasks:

-   a. System calls ChipP's Read function, passing in (dontCare, 1,     dontCare) -   b. ChipP returns M₁, from which System can extract f_(P) -   c. System stores f_(P) for future use

ChipP's key is updated as follows:

-   a. System calls ChipP's Random function; -   b. ChipP returns R_(P) to System; -   c. System calls ChipF's GetProgramKey function, passing in f_(P), n2     (the desired key to use) and the result from b; -   d. ChipF updates R_(F), then calculates and returns R_(F),     S_(fP(Kn2))[R_(P)|R_(F)|C₃]⊕K_(new), and     S_(fP(Kn2))[R_(F)|S_(fP(Kn2))[R_(P)|R_(F)|C₃]⊕K_(new)|C₃]; -   e. If the response from d is not 0, System calls ChipP's ReplaceKey     function, passing in n1 (the key to use in ChipP) and the response     from d; -   f. System checks response from ChipP. If the response is 1, then     ChipP's K_(n1) has been correctly updated to K_(new). If the     response is 0, ChipP's K_(n1) has not been updated.

The choice of n1 and n2 must be such that ChipP's K_(n1)=ChipF's f_(P)(K_(n2)).

5.5.4 GetProgramKey4—Indirect to Indirect

With the “old key=indirect, new key=indirect” form of GetProgramKey, to update P's key, the System must extract f from ChipP (so it can tell ChipF how to generate the correct key) by performing the following tasks:

-   a. System calls ChipP's Read function, passing in (dontCare, 1,     dontCare) -   b. ChipP returns M₁, from which System can extract f_(P) -   c. System stores f_(P) for future use

ChipP's key is updated as follows:

-   a. System calls ChipP's Random function; -   b. ChipP returns R_(P) to System; -   c. System calls ChipF's GetProgramKey function, passing in f_(P), n2     (the desired key to use) and the result from b; -   d. ChipF updates R_(F), then calculates and returns R_(F),     S_(fP(Kn2))[R_(P)|R_(F)|C₃]⊕f_(P)(K_(new)), and     S_(fP(Kn2))[R_(F)|S_(fP(Kn2))[R_(P)|R_(F)|C₃]⊕f_(P)(K_(new))|C₃]; -   e. If the response from d is not 0, System calls ChipP's ReplaceKey     function, passing in n1 (the key to use in ChipP) and the response     from d; -   f. System checks response from ChipP. If the response is 1, then     ChipP's K_(n1) has been correctly updated to f_(P)(K_(new)). If the     response is 0, ChipP's K_(n1) has not been updated.

The choice of n1 and n2 must be such that ChipP's K_(n1)=ChipF's f_(P)(K_(n2)).

5.5.5 Chicken and Egg

The Program Key protocol requires both ChipF and ChipP to know K_(old) (either directly or indirectly). Obviously both chips had to be programmed in some way with K_(old), and thus K_(old) can be thought of as an older K_(new):K_(old) can be placed in chips if another ChipF knows K_(older), and so on.

Although this process allows a chain of reprogramming of keys, with each stage secure, at some stage the very first key (K_(first)) must be placed in the chips. K_(first) is in fact programmed with the chip's microcode at the manufacturing test station as the last step in manufacturing test. K_(first) can be a manufacturing batch key, changed for each batch or for each customer etc., and can have as short a life as desired. Compromising K_(first) need not result in a complete compromise of the chain of Ks. This is especially true if K_(first) is indirectly stored in ChipPs (i.e. each ChipP holds an f and f(K_(first)) instead of K_(first) directly). One example is where K_(first) (the key stored in each chip after manufacture/test) is a batch key, and can be different per chip. K_(first) may advance to a ComCo specific K_(seond) etc. but still remain indirect. A direct form (e.g. K_(fianal)) only needs to go in if it is actually required at the end of the programming chain.

Depending on reprogramming requirements, K_(first) can be the same or different for all K_(n).

6 Memjet Forms of Protocols

Physical QA Chips are used in Memjet printer systems to store printer operating parameters as well as consumable parameters.

6.1 PRINTER_QA

A PRINTER_QA is stored within each print engine to perform two primary tasks:

-   -   storage and protection of operating parameters     -   a means of indirect read validation of other QA Chip data         vectors

Each PRINTER_QA contains the following keys:

TABLE 229 Keys in PrinterQA Key Contents Comments 0 Upgrade Key Used to upgrade the operating parameters. Should be indirect form of key (i.e. a different key for each PRINTER_QA) so that an indirect form of the write is required. 1 Consumable Read Used to indirectly read the data from Validation Key an CONSUMABLE_QA chip using indirect authenticated read protocol (Section 5.3.2 on page 606). 2 PrintEngineController When reading data from the Read Validation Key PRINTER_QA, the system can either rust the data, or must use this key to perform the authenticated read protocol (see Section 5.3 on page 604). 3-n (reserved) Currently unused. Could be used to provide a means to indirectly read additional print engine operating parameters ala K1, or provide additional Print Engine validation ala K2. Note that if multiple Print Engine controllers are used (e.g. a multiple SoPEC sytem), then multiple PrintEngineController Read Validation Keys are required. These keys can be stored within a single PRINTER_QA (e.g. in K₃ and beyond), or can be stored in separate PRINTER_QAs (for example each SoPEC (or group of SoPECs) has an individual PRINTER_QA).

The functions required in the PRINTER_QA are:

-   -   Random, ReplaceKey, to allow key programming & substitution     -   Read, to allow reads of data     -   Write, to allow updates of M₁₊ during manufacture     -   WriteAuth, to provide a means of updating the M₀ data (operating         parameters)     -   SetPermissions, to provide a means of updating write permissions     -   Test, to provide a means of checking if consumable reads are         valid     -   Translate, to provide a means of indirect reading of consumable         data         6.2 CONSUMABLE_QA

A CONSUMABLE_QA is stored with each consumable (e.g. ink cartridge) to perform two primary tasks:

-   -   storage of consumable related data     -   protection of consumable amount remaining

Each CONSUMABLE_QA contains the following keys:

TABLE 230 Keys in CONSUMABLE_QA Key Contents Comments 0 Upgrade Key Used to upgrade the consumable parameters. Should be stored as the indirect form of the key (i.e. a different key for each CONSUMABLE_QA) so that an indirect form of the write is required. 1 Consumable Read When reading data from the Validation Key CONSUMABLE_QA, the system can either trust the data, or must use this key to perform either the direct or indirect authenticated read protocol (see Section 5.3 on page 604). 2 (reserved) Currently unused. 3-n (reserved) Currently unused.

The functions required in the CONSUMABLE_QA are:

-   -   Random, ReplaceKey, to allow key programming & substitution     -   Read, to allow reads of data.     -   Write, to allow updates of M₁₊ during manufacture     -   WriteAuth, to provide a means of updating the M₀ data         (consumable remaining)     -   SetPermissions, to provide a means of updating write permissions         Authentication of Consumables         1 Introduction

Manufacturers of systems that require consumables (such as a laser printer that requires toner cartridges) have struggled with the problem of authenticating consumables, to varying levels of success. Most have resorted to specialized packaging that involves a patent. However this does not stop home refill operations or clone manufacture in countries with weak industrial property protection. The prevention of copying is important to prevent poorly manufactured substitute consumables from damaging the base system. For example, poorly filtered ink may clog print nozzles in an ink jet printer, causing the consumer to blame the system manufacturer and not admit the use of non-authorized consumables.

To solve the authentication problem, this document describes an QA Chip that contains authentication keys and circuitry specially designed to prevent copying. The chip is manufactured using the standard Flash memory manufacturing process, and is low cost enough to be included in consumables such as ink and toner cartridges. The implementation is approximately 1 mm² in a 0.25 micron flash process, and has an expected manufacturing cost of approximately 10 cents in 2003.

2 NSA

Once programmed, the QA Chips as described here are compliant with the NSA export guidelines since they do not constitute a strong encryption device. They can therefore be practically manufactured in the USA (and exported) or anywhere else in the world.

3 Nomenclature

The following symbolic nomenclature is used throughout this document:

TABLE 231 Summary of symbolic nomenclature Symbol Description F[X] Function F, taking a single parameter X F[X, Y] Function F, taking two parameters, X and Y X | Y X concatenated with Y X

Y Bitwise X AND Y X

Y Bitwise X OR Y (inclusive-OR) X ⊕ Y Bitwise X XOR Y (exclusive-OR)

X Bitwise NOT X (complement) X

Y X is assigned the value Y X

{Y, Z} The domain of assignment inputs to X is Y and Z X = Y X is equal to Y X ≠ Y X is not equal to Y

X Decrement X by 1 (floor 0)

X Increment X by 1 (modulo register length) Erase X Erase Flash memory register X SetBits[X, Y] Set the bits of the Flash memory register X based on Y Z

ShiftRight[X, Y] Shift register X right one bit position, taking input bit from Y and placing the output bit in Z 4 Pseudocode 4.1.1 Asynchronous

The following pseudocode:

-   -   var=expression         means the var signal or output is equal to the evaluation of the         expression.         4.1.2 Synchronous

The following pseudocode:

-   -   var←expression         means the var register is assigned the result of evaluating the         expression during this cycle.         4.1.3 Expression

Expressions are defined using the nomenclature in Table 231 above. Therefore:

-   -   var=(a=b)         is interpreted as the var signal is 1 if a is equal to b, and 0         otherwise.         4.2 Diagrams

Black is used to denote data, and red to denote 1-bit control-signal lines.

4.3 QA Chip Terminology

This document refers to QA Chips by their function in particular protocols:

-   -   For authenticated reads, ChipA is the QA Chip being         authenticated, and ChipT is the QA Chip that is trusted.     -   For replacement of keys, ChipP is the QA Chip being programmed         with the new key, and ChipF is the factory QA Chip that         generates the message to program the new key.     -   For upgrades of data in a QA Chip, ChipU is the QA Chip being         upgraded, and ChipS is the QA Chip that signs the upgrade value.

Any given physical QA Chip will contain functionality that allows it to operate as an entity in some number of these protocols.

Therefore, wherever the terms ChipA, ChipT, ChipP, ChipF, ChipU and ChipS are used in this document, they are referring to logical entities involved in an authentication protocol as defined in subsequent sections.

Physical QA Chips are referred to by their location. For example, each ink cartridge may contain a QA Chip referred to as an INK_QA, with all INK_QA chips being on the same physical bus. In the same way, the QA Chip inside a printer is referred to as PRINTER_QA, and will be on a separate bus to the INK_QA chips.

5 Concepts and Terms

This chapter provides a background to the problem of authenticating consumables. For more in-depth introductory texts, see [12], [78], and [56].

5.1 Basic Terms

A message, denoted by M, is plaintext. The process of transforming M into ciphertext C, where the substance of M is hidden, is called encryption. The process of transforming C back into M is called decryption. Referring to the encryption function as E, and the decryption function as D, we have the following identities: E[M]=C D[C]=M

Therefore the following identity is true: D[E[M]]=M 5.2 Symmetric Cryptography

A symmetric encryption algorithm is one where:

-   -   the encryption function E relies on key K₁,     -   the decryption function D relies on key K₂,     -   K₂ can be derived from K₁, and     -   K₁ can be derived from K₂.

In most symmetric algorithms, K₁ equals K₂. However, even if K₁ does not equal K₂, given that one key can be derived from the other, a single key K can suffice for the mathematical definition. Thus: E_(K)[M]=C D_(K)[C]=M

The security of these algorithms rests very much in the key K. Knowledge of K allows anyone to encrypt or decrypt. Consequently K must remain a secret for the duration of the value of M. For example, M may be a wartime message “My current position is grid position 123-456”. Once the war is over the value of M is greatly reduced, and if K is made public, the knowledge of the combat unit's position may be of no relevance whatsoever. Of course if it is politically sensitive for the combat unit's position to be known even after the war, K may have to remain secret for a very long time.

An enormous variety of symmetric algorithms exist, from the textbooks of ancient history through to sophisticated modern algorithms. Many of these are insecure, in that modern cryptanalysis techniques (see Section 5.7 on page 646) can successfully attack the algorithm to the extent that K can be derived.

The security of the particular symmetric algorithm is a function of two things: the strength of the algorithm and the length of the key [78].

The strength of an algorithm is difficult to quantify, relying on its resistance to cryptographic attacks (see Section 5.7 on page 646). In addition, the longer that an algorithm has remained in the public eye, and yet remained unbroken in the midst of intense scrutiny, the more secure the algorithm is likely to be. By contrast, a secret algorithm that has not been scrutinized by cryptographic experts is unlikely to be secure.

Even if the algorithm is “perfectly” strong (the only way to break it is to try every key—see Section 5.7.1.5 on page 647), eventually the right key will be found. However, the more keys there are, the more keys have to be tried. If there are N keys, it will take a maximum of N tries. If the key is N bits long, it will take a maximum of 2^(N) tries, with a 50% chance of finding the key after only half the attempts (2^(N−1)). The longer N becomes, the longer it will take to find the key, and hence the more secure it is. What makes a good key length depends on the value of the secret and the time for which the secret must remain secret as well as available computing resources.

In 1996, an ad hoc group of world-renowned cryptographers and computer scientists released a report [9] describing minimal key lengths for symmetric ciphers to provide adequate commercial security. They suggest an absolute minimum key length of 90 bits in order to protect data for 20 years, and stress that increasingly, as cryptosystems succumb to smarter attacks than brute-force key search, even more bits may be required to account for future surprises in cryptanalysis techniques.

We will ignore most historical symmetric algorithms on the grounds that they are insecure, especially given modern computing technology. Instead, we will discuss the following algorithms:

-   -   DES     -   Blowfish     -   RC5     -   IDEA         5.2.1 DES

DES (Data Encryption Standard) [26] is a US and international standard, where the same key is used to encrypt and decrypt. The key length is 56 bits. It has been implemented in hardware and software, although the original design was for hardware only. The original algorithm used in DES was patented in 1976 (U.S. Pat. No. 3,962,539) and has since expired.

During the design of DES, the NSA (National Security Agency) provided secret S-boxes to perform the key-dependent nonlinear transformations of the data block. After differential cryptanalysis was discovered outside the NSA, it was revealed that the DES S-boxes were specifically designed to be resistant to differential cryptanalysis.

As described in [95], using 1993 technology, a 56-bit DES key can be recovered by a custom-designed $1 million machine performing a brute force attack in only 35 minutes. For $10 million, the key can be recovered in only 3.5 minutes. DES is clearly not secure now, and will become less so in the future.

A variant of DES, called triple-DES is more secure, but requires 3 keys: K₁, K₂, and K₃. The keys are used in the following manner: E_(K3) [D_(K2)[E_(K1)[M]]]=C D_(K3)[E_(K2)[D_(K1[C]]]=M)

The main advantage of triple-DES is that existing DES implementations can be used to give more security than single key DES. Specifically, triple-DES gives protection of equivalent key length of 112 bits [78]. Triple-DES does not give the equivalent protection of a 168-bit key (3×56) as one might naively expect.

Equipment that performs triple-DES decoding and/or encoding cannot be exported from the United States.

5.2.2 Blowfish

Blowfish is a symmetric block cipher first presented by Schneier in 1994 [76]. It takes a variable length key, from 32 bits to 448 bits, is unpatented, and is both license and royalty free. In addition, it is much faster than DES.

The Blowfish algorithm consists of two parts: a key-expansion part and a data-encryption part. Key expansion converts a key of at most 448 bits into several subkey arrays totaling 4168 bytes. Data encryption occurs via a 16-round Feistel network. All operations are XORs and additions on 32-bit words, with four index array lookups per round.

It should be noted that decryption is the same as encryption except that the subkey arrays are used in the reverse order. Complexity of implementation is therefore reduced compared to other algorithms that do not have such symmetry.

[77] describes the published attacks which have been mounted on Blowfish, although the algorithm remains secure as of February 1998 [79]. The major finding with these attacks has been the discovery of certain weak keys. These weak keys can be tested for during key generation. For more information, refer to [77] and [79].

5.2.3 RC5

Designed by Ron Rivest in 1995, RC5 [74] has a variable block size, key size, and number of rounds. Typically, however, it uses a 64-bit block size and a 128-bit key.

The RC5 algorithm consists of two parts: a key-expansion part and a data-encryption part. Key expansion converts a key into 2r+2 subkeys (where r=the number of rounds), each subkey being w bits. For a 64-bit blocksize with 16 rounds (w=32, r=16), the subkey arrays total 136 bytes. Data encryption uses addition mod 2^(W), XOR and bitwise rotation.

An initial examination by Kaliski and Yin [43] suggested that standard linear and differential cryptanalysis appeared impractical for the 64-bit blocksize version of the algorithm. Their differential attacks on 9 and 12 round RC5 require 2⁴⁵ and 2⁶² chosen plaintexts respectively, while the linear attacks on 4, 5, and 6 round RC5 requires 2³⁷, 2⁴⁷ and 2⁵⁷ known plaintexts). These two attacks are independent of key size.

More recently however, Knudsen and Meier [47] described a new type of differential attack on RC5 that improved the earlier results by a factor of 128, showing that RC5 has certain weak keys.

RC5 is protected by multiple patents owned by RSA Laboratories. A license must be obtained to use it.

5.2.4 IDEA

Developed in 1990 by Lai and Massey [53], the first incarnation of the IDEA cipher was called PES. After differential cryptanalysis was discovered by Biham and Shamir in 1991, the algorithm was strengthened, with the result being published in 1992 as IDEA [52].

IDEA uses 128-bit keys to operate on 64-bit plaintext blocks. The same algorithm is used for encryption and decryption. It is generally regarded as the most secure block algorithm available today [78][78].

The biggest drawback of IDEA is the fact that it is patented (U.S. Pat. No. 5,214,703, issued in 1993), and a license must be obtained from Ascom Tech AG (Bern) to use it.

5.3 Asymmetric Cryptography

An asymmetric encryption algorithm is one where:

-   -   the encryption function E relies on key K₁,     -   the decryption function D relies on key K₂,     -   K₂ cannot be derived from K₁ in a reasonable amount of time, and     -   K₁ cannot be derived from K₂ in a reasonable amount of time.

Thus: E_(K1)[M]=C D_(K2)[C]=M

These algorithms are also called public-key because one key K₁ can be made public. Thus anyone can encrypt a message (using K₁) but only the person with the corresponding decryption key (K₂) can decrypt and thus read the message.

In most cases, the following identity also holds: E_(K2)[M]=C D_(K1)[C]=M

This identity is very important because it implies that anyone with the public key K₁ can see M and know that it came from the owner of K₂. No-one else could have generated C because to do so would imply knowledge of K₂. This gives rise to a different application, unrelated to encryption—digital signatures.

The property of not being able to derive K₁ from K₂ and vice versa in a reasonable time is of course clouded by the concept of reasonable time. What has been demonstrated time after time, is that a calculation that was thought to require a long time has been made possible by the introduction of faster computers, new algorithms etc. The security of asymmetric algorithms is based on the difficulty of one of two problems: factoring large numbers (more specifically large numbers that are the product of two large primes), and the difficulty of calculating discrete logarithms in a finite field. Factoring large numbers is conjectured to be a hard problem given today's understanding of mathematics. The problem however, is that factoring is getting easier much faster than anticipated. Ron Rivest in 1977 said that factoring a 125-digit number would take 40 quadrillion years [30]. In 1994 a 129-digit number was factored (3]. According to Schneier, you need a 1024-bit number to get the level of security today that you got from a 512-bit number in the 1980s [78]. If the key is to last for some years then 1024 bits may not even be enough. Rivest revised his key length estimates in 1990: he suggests 1628 bits for high security lasting until 2005, and 1884 bits for high security lasting until 2015 [69]. Schneier suggests 2048 bits are required in order to protect against corporations and governments until 2015 [80].

Public key cryptography was invented in 1976 by Diffie and Hellman [15][15], and independently by Merkle [57]. Although Diffie, Hellman and Merkle patented the concepts (U.S. Pat. Nos. 4,200,770 and 4,218,582), these patents expired in 1997.

A number of public key cryptographic algorithms exist. Most are impractical to implement, and many generate a very large C for a given M or require enormous keys. Still others, while secure, are far too slow to be practical for several years. Because of this, many public key systems are hybrid—a public key mechanism is used to transmit a symmetric session key, and then the session key is used for the actual messages.

All of the algorithms have a problem in terms of key selection. A random number is simply not secure enough. The two large primes p and q must be chosen carefully—there are certain weak combinations that can be factored more easily (some of the weak keys can be tested for). But nonetheless, key selection is not a simple matter of randomly selecting 1024 bits for example. Consequently the key selection process must also be secure.

Of the practical algorithms in use under public scrutiny, the following are discussed:

-   -   RSA     -   DSA     -   EIGamal         5.3.1 RSA

The RSA cryptosystem [75], named after Rivest, Shamir, and Adleman, is the most widely used public key cryptosystem, and is a de facto standard in much of the world [78].

The security of RSA depends on the conjectured difficulty of factoring large numbers that are the product of two primes (p and q). There are a number of restrictions on the generation of p and q. They should both be large, with a similar number of bits, yet not be close to one another (otherwise p≡q≡√pq). In addition, many authors have suggested that p and q should be strong primes [56]. The Hellman-Bach patent (U.S. Pat. No. 4,633,036) covers a method for generating strong RSA primes p and q such that n=pq and factoring n is believed to be computationally infeasible. The RSA algorithm patent was issued in 1983 (U.S. Pat. No. 4,405,829). The patent expires on Sep. 20, 2000.

5.3.2 DSA

DSA (Digital Signature Algorithm) is an algorithm designed as part of the Digital Signature Standard (DSS) [29]. As defined, it cannot be used for generalized encryption. In addition, compared to RSA, DSA is 10 to 40 times slower for signature verification [40]. DSA explicitly uses the SHA-1 hashing algorithm (see Section 5.5.3.3 on page 640).

DSA key generation relies on finding two primes p and q such that q divides p−1. According to Schneier [78], a 1024-bit p value is required for long term DSA security. However the DSA standard [29] does not permit values of p larger than 1024 bits (p must also be a multiple of 64 bits).

The US Government owns the DSA algorithm and has at least one relevant patent (U.S. Pat. No. 5,231,688 granted in 1993). However, according to NIST [61]:

-   -   “The DSA patent and any foreign counterparts that may issue are         available for use without any written permission from or any         payment of royalties to the U.S. government.”

In a much stronger declaration, NIST states in the same document [61] that DSA does not infringe third party's rights:

-   -   “NIST reviewed all of the asserted patents and concluded that         none of them would be infringed by DSS. Extra protection will be         written into the PK1 pilot project that will prevent an         organization or individual from suing anyone except the         government for patent infringement during the course of the         project.”

It must however, be noted that the Schnorr authentication algorithm [81] (U.S. Pat. No. 4,995,082) patent holder claims that DSA infringes his patent. The Schnorr patent is not due to expire until 2008.

5.3.3 EIGamal

The EIGamal scheme [22][22] is used for both encryption and digital signatures. The security is based on the conjectured difficulty of calculating discrete logarithms in a finite field.

Key selection involves the selection of a prime p, and two random numbers g and x such that both g and x are less than p. Then calculate y=gx mod p. The public key is y, g, and p. The private key is x.

EIGamal is unpatented. Although it uses the patented Diffie-Hellman public key algorithm [15][15], those patents expired in 1997. EIGamal public key encryption and digital signatures can now be safely used without infringing third party patents.

5.4 Cryptographic Challenge-Response Protocols and Zero Knowledge Proofs

The general principle of a challenge-response protocol is to provide identity authentication. The simplest form of challenge-response takes the form of a secret password. A asks B for the secret password, and if B responds with the correct password, A declares B authentic.

There are three main problems with this kind of simplistic protocol. Firstly, once B has responded with the password, any observer C will know what the password is. Secondly, A must know the password in order to verify it. Thirdly, if C impersonates A, then B will give the password to C (thinking C was A), thus compromising the password.

Using a copyright text (such as a haiku) as the password is not sufficient, because we are assuming that anyone is able to copy the password (for example in a country where intellectual property is not respected).

The idea of cryptographic challenge-response protocols is that one entity (the claimant) proves its identity to another (the verifier) by demonstrating knowledge of a secret known to be associated with that entity, without revealing the secret itself to the verifier during the protocol [56]. In the generalized case of cryptographic challenge-response protocols, with some schemes the verifier knows the secret, while in others the secret is not even known by the verifier. A good overview of these protocols can be found in [25], [78], and [56].

Since this documentation specifically concerns Authentication, the actual cryptographic challenge-response protocols used for authentication are detailed in the appropriate sections. However the concept of Zero Knowledge Proofs bears mentioning here.

The Zero Knowledge Proof protocol, first described by Feige, Fiat and Shamir in [24] is extensively used in Smart Cards for the purpose of authentication [34][34][34]. The protocol's effectiveness is based on the assumption that it is computationally infeasible to compute square roots modulo a large composite integer with unknown factorization. This is provably equivalent to the assumption that factoring large integers is difficult.

It should be noted that there is no need for the claimant to have significant computing power. Smart cards implement this kind of authentication using only a few modulo multiplications [34][34].

Finally, it should be noted that the Zero Knowledge Proof protocol is patented [82] (U.S. Pat. No. 4,748,668, issued May 31, 1988).

5.5 One-Way Functions

A one-way function F operates on an input X, and returns F[X] such that X cannot be determined from F[X]. When there is no restriction on the format of X, and F[X] contains fewer bits than X, then collisions must exist. A collision is defined as two different X input values producing the same F[X] value—i.e. X₁ and X₂ exist such that X₁≠X₂ yet F[X₁]=F[X₂].

When X contains more bits than F[X], the input must be compressed in some way to create the output. In many cases, X is broken into blocks of a particular size, and compressed over a number of rounds, with the output of one round being the input to the next. The output of the hash function is the last output once X has been consumed. A pseudo-collision of the compression function CF is defined as two different initial values V₁ and V₂ and two inputs X₁ and X₂ (possibly identical) are given such that CF(V₁, X₁)=CF(V₂, X₂). Note that the existence of a pseudo-collision does not mean that it is easy to compute an X₂ for a given X₁.

We are only interested in one-way functions that are fast to compute. In addition, we are only interested in deterministic one-way functions that are repeatable in different implementations.

Consider an example F where F[X] is the time between calls to F. For a given F(X] X cannot be determined because X is not even used by F. However the output from F will be different for different implementations. This kind of F is therefore not of interest.

In the scope of this document, we are interested in the following forms of one-way functions:

-   -   Encryption using an unknown key     -   Random number sequences     -   Hash Functions     -   Message Authentication Codes         5.5.1 Encryption Using an Unknown Key

When a message is encrypted using an unknown key K, the encryption function E is effectively one-way. Without the key, it is computationally infeasible to obtain M from EK[M] without K. An encryption function is only one-way for as long as the key remains hidden.

An encryption algorithm does not create collisions, since E creates EK[M] such that it is possible to reconstruct M using function D. Consequently F[X] contains at least as many bits as X (no information is lost) if the one-way function F is E.

Symmetric encryption algorithms (see Section 5.2 on page 629) have the advantage over asymmetric algorithms (see Section 5.3 on page 632) for producing one-way functions based on encryption for the following reasons:

-   -   The key for a given strength encryption algorithm is shorter for         a symmetric algorithm than an asymmetric algorithm     -   Symmetric algorithms are faster to compute and require less         software or silicon

Note however, that the selection of a good key depends on the encryption algorithm chosen. Certain keys are not strong for particular encryption algorithms, so any key needs to be tested for strength. The more tests that need to be performed for key selection, the less likely the key will remain hidden.

5.5.2 Random Number Sequences

Consider a random number sequence R₀, R₁, . . . , R_(i), R_(i+1). We define the one-way function F such that F[X] returns the X^(th) random number in the random sequence. However we must ensure that F[X] is repeatable for a given X on different implementations. The random number sequence therefore cannot be truly random. Instead, it must be pseudo-random, with the generator making use of a specific seed.

There are a large number of issues concerned with defining good random number generators. Knuth, in [48] describes what makes a generator “good” (including statistical tests), and the general problems associated with constructing them. Moreau gives a high level survey of the current state of the field in [60].

The majority of random number generators produce the i^(th) random number from the i−1 ^(th) state—the only way to determine the i^(th) number is to iterate from the 0^(th) number to the i^(th). If i is large, it may not be practical to wait for i iterations.

However there is a type of random number generator that does allow random access. In [10], Blum, Blum and Shub define the ideal generator as follows: “ . . . we would like a pseudo-random sequence generator to quickly produce, from short seeds, long sequences (of bits) that appear in every way to be generated by successive flips of a fair coin”. They defined the x² mod n generator [10], more commonly referred to as the BBS generator. They showed that given certain assumptions upon which modern cryptography relies, a BBS generator passes extremely stringent statistical tests.

The BBS generator relies on selecting n which is a Blum integer (n=pq where p and q are large prime numbers, p≠q, p mod 4=3, and q mod 4=3). The initial state of the generator is given by x₀ where x₀=x² mod n, and x is a random integer relatively prime to n. The i^(th) pseudo-random bit is the least significant bit of x_(i) where: x _(i) =x _(i−1) ² mod n

As an extra property, knowledge of p and q allows a direct calculation of the i^(th) number in the sequence as follows: x _(i) =x ₀ ^(y) mod n where y=2^(i) mod((p−1)(q−1))

Without knowledge of p and q, the generator must iterate (the security of calculation relies on the conjectured difficulty of factoring large numbers).

When first defined, the primary problem with the BBS generator was the amount of work required for a single output bit. The algorithm was considered too slow for most applications. However the advent of Montgomery reduction arithmetic [58] has given rise to more practical implementations, such as [59]. In addition, Vazirani and Vazirani have shown in [93] that depending on the size of n, more bits can safely be taken from x_(i) without compromising the security of the generator.

Assuming we only take 1 bit per x_(i), N bits (and hence N iterations of the bit generator function) are needed in order to generate an N-bit random number. To the outside observer, given a particular set of bits, there is no way to determine the next bit other than a 50/50 probability. If the x, p and q are hidden, they act as a key, and it is computationally infeasible to take an output bit stream and compute x, p, and q. It is also computationally infeasible to determine the value of i used to generate a given set of pseudo-random bits. This last feature makes the generator one-way.

Different values of i can produce identical bit sequences of a given length (e.g. 32 bits of random bits). Even if x, p and q are known, for a given F[i], i can only be derived as a set of possibilities, not as a certain value (of course if the domain of i is known, then the set of possibilities is reduced further).

However, there are problems in selecting a good p and q, and a good seed x. In particular, Ritter in [68] describes a problem in selecting x. The nature of the problem is that a BBS generator does not create a single cycle of known length. Instead, it creates cycles of various lengths, including degenerate (zero-length) cycles. Thus a BBS generator cannot be initialized with a random state—it might be on a short cycle. Specific algorithms exist in section 9 of [10] to determine the length of the period for a given seed given certain strenuous conditions for n.

5.5.3 Hash Functions

Special one-way functions, known as Hash functions, map arbitrary length messages to fixed-length hash values. Hash functions are referred to as H[M]. Since the input is of arbitrary length, a hash function has a compression component in order to produce a fixed length output. Hash functions also have an obfuscation component in order to make it difficult to find collisions and to determine information about M from H[M].

Because collisions do exist, most applications require that the hash algorithm is preimage resistant, in that for a given X₁ it is difficult to find X₂ such that H[X₁]=H[X₂]. In addition, most applications also require the hash algorithm to be collision resistant (i.e. it should be hard to find two messages X₁ and X₂ such that H[X₁]=H[X₂]). However, as described in [20], it is an open problem whether a collision-resistant hash function, in the ideal sense, can exist at all.

The primary application for hash functions is in the reduction of an input message into a digital “fingerprint” before the application of a digital signature algorithm. One problem of collisions with digital signatures can be seen in the following example.

-   -   A has a long message M₁ that says “I owe B $10”. A signs H[M₁]         using his private key. B, being greedy, then searches for a         collision message M₂ where H[M₂]=H[M₁] but where M₂ is favorable         to B, for example “I owe B $1 million”. Clearly it is in A's         interest to ensure that it is difficult to find such an M₂.

Examples of collision resistant one-way hash functions are SHA-1 [28], MD5 [73] and RIPEMD-160 [66], all derived from MD4 [70][70].

5.5.3.1 MD4

Ron Rivest introduced MD4 [70][70] in 1990. It is only mentioned here because all other one-way hash functions are derived in some way from MD4.

MD4 is now considered completely broken [18][18] in that collisions can be calculated instead of searched for. In the example above, B could trivially generate a substitute message M₂ with the same hash value as the original message M₁.

5.5.3.2 MD5

Ron Rivest introduced MD5 [73] in 1991 as a more secure MD4. Like MD4, MD5 produces a 128-bit hash value. MD5 is not patented [80].

Dobbertin describes the status of MD5 after recent attacks [20]. He describes how pseudo-collisions have been found in MD5, indicating a weakness in the compression function, and more recently, collisions have been found. This means that MD5 should not be used for compression in digital signature schemes where the existence of collisions may have dire consequences. However MD5 can still be used as a one-way function. In addition, the HMAC-MD5 construct (see Section 5.5.4.1 on page 643) is not affected by these recent attacks.

5.5.3.3 SHA-1

SHA-1 [28] is very similar to MD5, but has a 160-bit hash value (MD5 only has 128 bits of hash value). SHA-1 was designed and introduced by the NIST and NSA for use in the Digital Signature Standard (DSS). The original published description was called SHA [27], but very soon afterwards, was revised to become SHA-1 [28], supposedly to correct a security flaw in SHA (although the NSA has not released the mathematical reasoning behind the change).

There are no known cryptographic attacks against SHA-1 [78]. It is also more resistant to brute force attacks than MD4 or MD5 simply because of the longer hash result.

The U.S. Government owns the SHA-1 and DSA algorithms (a digital signature authentication algorithm defined as part of DSS [29]) and has at least one relevant patent (U.S. Pat. No. 5,231,688 granted in 1993). However, according to NIST [61]:

-   -   “The DSA patent and any foreign counterparts that may issue are         available for use without any written permission from or any         payment of royalties to the U.S. government.”

In a much stronger declaration, NIST states in the same document [61] that DSA and SHA-1 do not infringe third party's rights:

-   -   “NIST reviewed all of the asserted patents and concluded that         none of them would be infringed by DSS. Extra protection will be         written into the PK1 pilot project that will prevent an         organization or individual from suing anyone except the         government for patent infringement during the course of the         project.”

It must however, be noted that the Schnorr authentication algorithm [81] (U.S. Pat. No. 4,995,082) patent holder claims that DSA infringes his patent. The Schnorr patent is not due to 20 expire until 2008. Fortunately this does not affect SHA-1.

5.5.3.4 RIPEMD-160

RIPEMD-160 [66] is a hash function derived from its predecessor RIPEMD [11] (developed for the European Community's RIPE project in 1992). As its name suggests, RIPEMD-160 produces a 160-bit hash result. Tuned for software implementations on 32-bit architectures, RIPEMD-160 is intended to provide a high level of security for 10 years or more.

Although there have been no successful attacks on RIPEMD-160, it is comparatively new and has not been extensively cryptanalyzed. The original RIPEMD algorithm [11] was specifically designed to resist known cryptographic attacks on MD4. The recent attacks on MD5 (detailed in [20]) showed similar weaknesses in the RIPEMD 128-bit hash function. Although the attacks showed only theoretical weaknesses, Dobbertin, Preneel and Bosselaers further strengthened RIPEMD into a new algorithm RIPEMD-160.

RIPEMD-160 is in the public domain, and requires no licensing or royalty payments.

5.5.4 Message Authentication Codes

The problem of message authentication can be summed up as follows:

-   -   How can A be sure that a message supposedly from B is in fact         from B?

Message authentication is different from entity authentication (described in the section on cryptographic challenge-response protocols). With entity authentication, one entity (the claimant) proves its identity to another (the verifier). With message authentication, we are concerned with making sure that a given message is from who we think it is from i.e. it has not been tampered with en route from the source to its destination. While this section has a brief overview of message authentication, a more detailed survey can be found in [88].

A one-way hash function is not sufficient protection for a message. Hash functions such as MD5 rely on generating a hash value that is representative of the original input, and the original input cannot be derived from the hash value. A simple attack by E, who is in-between A and B, is to intercept the message from B, and substitute his own. Even if A also sends a hash of the original message, E can simply substitute the hash of his new message. Using a one-way hash function alone, A has no way of knowing that B's message has been changed.

One solution to the problem of message authentication is the Message Authentication Code, or MAC.

When B sends message M, it also sends MAC(M] so that the receiver will know that M is actually from B. For this to be possible, only B must be able to produce a MAC of M, and in addition, A should be able to verify M against MAC[M]. Notice that this is different from encryption of M−MACs are useful when M does not have to be secret.

The simplest method of constructing a MAC from a hash function is to encrypt the hash value with a symmetric algorithm:

-   1. Hash the input message H[M] -   2. Encrypt the hash E_(K)[H[M]]

This is more secure than first encrypting the message and then hashing the encrypted message. Any symmetric or asymmetric cryptographic function can be used, with the appropriate advantages and disadvantage of each type described in Section 5.2 on page 629 and Section 5.3 on page 632.

However, there are advantages to using a key-dependent one-way hash function instead of techniques that use encryption (such as that shown above):

-   -   Speed, because one-way hash functions in general work much         faster than encryption;     -   Message size, because EK[M] is at least the same size as M,         while H[M] is a fixed size (usually considerably smaller than         M);     -   Hardware/software requirements—keyed one-way hash functions are         typically far less complex than their encryption-based         counterparts; and     -   One-way hash function implementations are not considered to be         encryption or decryption devices and therefore are not subject         to US export controls.

It should be noted that hash functions were never originally designed to contain a key or to support message authentication. As a result, some ad hoc methods of using hash functions to perform message authentication, including various functions that concatenate messages with secret prefixes, suffixes, or both have been proposed [56][56]. Most of these ad hoc methods have been successfully attacked by sophisticated means [42][42][42]. Additional MACs have been suggested based on XOR schemes [8] and Toeplitz matrices [49] (including the special case of LFSR-based (Linear Feed Shift Register) constructions).

5.5.4.1 HMAC

The HMAC construction [6][6] in particular is gaining acceptance as a solution for Internet message authentication security protocols. The HMAC construction acts as a wrapper, using the underlying hash function in a black-box way. Replacement of the hash function is straightforward if desired due to security or performance reasons. However, the major advantage of the HMAC construct is that it can be proven secure provided the underlying hash function has some reasonable cryptographic strengths—that is, HMAC's strengths are directly connected to the strength of the hash function [6].

Since the HMAC construct is a wrapper, any iterative hash function can be used in an HMAC. Examples include HMAC-MD5, HMAC-SHA1, HMAC-RIPEMD160 etc.

Given the following definitions:

-   -   H=the hash function (e.g. MD5 or SHA-1)     -   n=number of bits output from H (e.g. 160 for SHA-1, 128 bits for         MD5)     -   M=the data to which the MAC function is to be applied     -   K=the secret key shared by the two parties     -   ipad=0x36 repeated 64 times     -   opad=0x5C repeated 64 times

The HMAC algorithm is as follows:

-   1. Extend K to 64 bytes by appending 0x00 bytes to the end of K -   2. XOR the 64 byte string created in (1) with ipad -   3. append data stream M to the 64 byte string created in (2) -   4. Apply H to the stream generated in (3) -   5. XOR the 64 byte string created in (1) with opad -   6. Append the H result from (4) to the 64 byte string resulting from     (5) -   7. Apply H to the output of (6) and output the result

Thus: HMAC[M]=H[(K ⊕ opad)|H[(K ⊕ ipad)|M]]

The recommended key length is at least n bits, although it should not be longer than 64 bytes (the length of the hashing block). A key longer than n bits does not add to the security of the function.

HMAC optionally allows truncation of the final output e.g. truncation to 128 bits from 160 bits.

The HMAC designers' Request for Comments [51] was issued in 1997, one year after the algorithm was first introduced. The designers claimed that the strongest known attack against HMAC is based on the frequency of collisions for the hash function H (see Section 14.10 on page 700), and is totally impractical for minimally reasonable hash functions:

-   -   As an example, if we consider a hash function like MD5 where the         output length is 128 bits, the attacker needs to acquire the         correct message authentication tags computed (with the same         secret key K) on about 2⁶⁴ known plaintexts. This would require         the processing of at least 2⁶⁴ blocks under H, an impossible         task in any realistic scenario (for a block length of 64 bytes         this would take 250,000 years in a continuous 1 Gbps link, and         without changing the secret key K all this time). This attack         could become realistic only if serious flaws in the collision         behavior of the function H are discovered (e.g. Collisions found         after 2³⁰ messages). Such a discovery would determine the         immediate replacement of function H (the effects of such a         failure would be far more severe for the traditional uses of H         in the context of digital signatures, public key certificates         etc).

Of course, if a 160-bit hash function is used, then 2⁶⁴ should be replaced with 2⁸⁰.

This should be contrasted with a regular collision attack on cryptographic hash functions where no secret key is involved and 2⁶⁴ off-line parallelizable operations suffice to find collisions.

More recently, HMAC protocols with replay prevention components [62] have been defined in order to prevent the capture and replay of any M, HMAC[M] combination within a given time period.

Finally, it should be noted that HMAC is in the public domain [50], and incurs no licensing fees. There are no known patents infringed by HMAC.

5.6 Random Numbers and Time Varying Messages

The use of a random number generator as a one-way function has already been examined. However, random number generator theory is very much intertwined with cryptography, security, and authentication.

There are a large number of issues concerned with defining good random number generators. Knuth, in [48] describes what makes a generator good (including statistical tests), and the general problems associated with constructing them. Moreau gives a high level survey of the current state of the field in [60].

One of the uses for random numbers is to ensure that messages vary over time. Consider a system where A encrypts commands and sends them to B. If the encryption algorithm produces the same output for a given input, an attacker could simply record the messages and play them back to fool B. There is no need for the attacker to crack the encryption mechanism other than to know which message to play to B (while pretending to be A). Consequently messages often include a random number and a time stamp to ensure that the message (and hence its encrypted counterpart) varies each time.

Random number generators are also often used to generate keys. Although Klapper has recently shown [45] that a family of secure feedback registers for the purposes of building key-streams does exist, he does not give any practical construction. It is therefore best to say at the moment that all generators are insecure for this purpose. For example, the Berlekamp-Massey algorithm [54], is a classic attack on an LFSR random number generator. If the LFSR is of length n, then only 2n bits of the sequence suffice to determine the LFSR, compromising the key generator.

If, however, the only role of the random number generator is to make sure that messages vary over time, the security of the generator and seed is not as important as it is for session key generation. If however, the random number seed generator is compromised, and an attacker is able to calculate future “random” numbers, it can leave some protocols open to attack. Any new protocol should be examined with respect to this situation.

The actual type of random number generator required will depend upon the implementation and the purposes for which the generator is used. Generators include Blum, Blum, and Shub [10], stream ciphers such as RC4 by Ron Rivest [71], hash functions such as SHA-1 [28] and RIPEMD-160 [66], and traditional generators such LFSRs (Linear Feedback Shift Registers) [48] and their more recent counterpart FCSRs (Feedback with Carry Shift Registers) [44].

5.7 Attacks

This section describes the various types of attacks that can be undertaken to break an authentication cryptosystem. The attacks are grouped into physical and logical attacks.

Logical attacks work on the protocols or algorithms rather than their physical implementation, and attempt to do one of three things:

-   -   Bypass the authentication process altogether     -   Obtain the secret key by force or deduction, so that any         question can be answered     -   Find enough about the nature of the authenticating questions and         answers in order to, without the key, give the right answer to         each question.

Regardless of the algorithms and protocol used by a security chip, the circuitry of the authentication part of the chip can come under physical attack. Physical attacks come in four main ways, although the form of the attack can vary:

-   -   Bypassing the security chip altogether     -   Physical examination of the chip while in operation (destructive         and non-destructive)     -   Physical decomposition of chip     -   Physical alteration of chip

The attack styles and the forms they take are detailed below.

This section does not suggest solutions to these attacks. It merely describes each attack type. The examination is restricted to the context of an authentication chip (as opposed to some other kind of system, such as Internet authentication) attached to some System.

5.7.1 Logical Attacks

These attacks are those which do not depend on the physical implementation of the cryptosystem. They work against the protocols and the security of the algorithms and random number generators.

5.7.1.1 Ciphertext Only Attack

This is where an attacker has one or more encrypted messages, all encrypted using the same algorithm. The aim of the attacker is to obtain the plaintext messages from the encrypted messages. Ideally, the key can be recovered so that all messages in the future can also be recovered.

5.7.1.2 Known Plaintext Attack

This is where an attacker has both the plaintext and the encrypted form of the plaintext. In the case of an authentication chip, a known-plaintext attack is one where the attacker can see the data flow between the system and the authentication chip. The inputs and outputs are observed (not chosen by the attacker), and can be analyzed for weaknesses (such as birthday attacks or by a search for differentially interesting input/output pairs).

A known plaintext attack can be carried out by connecting a logic analyzer to the connection between the system and the authentication chip.

5.7.1.3 Chosen Plaintext Attacks

A chosen plaintext attack describes one where a cryptanalyst has the ability to send any chosen message to the cryptosystem, and observe the response. If the cryptanalyst knows the algorithm, there may be a relationship between inputs and outputs that can be exploited by feeding a specific output to the input of another function.

The chosen plaintext attack is much stronger than the known plaintext attack since the attacker can choose the messages rather than simply observe the data flow.

On a system using an embedded authentication chip, it is generally very difficult to prevent chosen plaintext attacks since the cryptanalyst can logically pretend he/she is the system, and thus send any chosen bit-pattern streams to the authentication chip.

5.7.1.4 Adaptive Chosen Plaintext Attacks

This type of attack is similar to the chosen plaintext attacks except that the attacker has the added ability to modify subsequent chosen plaintexts based upon the results of previous experiments. This is certainly the case with any system/authentication chip scenario described for consumables such as photocopiers and toner cartridges, especially since both systems and consumables are made available to the public.

5.7.1.5 Brute Force Attack

A guaranteed way to break any key-based cryptosystem algorithm is simply to try every key.

Eventually the right one will be found. This is known as a brute force attack. However, the more key possibilities there are, the more keys must be tried, and hence the longer it takes (on average) to find the right one. If there are N keys, it will take a maximum of N tries. If the key is N bits long, it will take a maximum of 2^(N) tries, with a 50% chance of finding the key after only half the attempts (2^(N−1)). The longer N becomes, the longer it will take to find the key, and hence the more secure the key is. Of course, an attack may guess the key on the first try, but this is more unlikely the longer the key is.

Consider a key length of 56 bits. In the worst case, all 2⁵⁶ tests (7.2×10¹⁶ tests) must be made to find the key. In 1977, Diffie and Hellman described a specialized machine for cracking DES, consisting of one million processors, each capable of running one million tests per second [17]. Such a machine would take 20 hours to break any DES code.

Consider a key length of 128 bits. In the worst case, all 2¹²⁸ tests (3.4×10³⁸ tests) must be made to find the key. This would take ten billion years on an array of a trillion processors each running 1 billion tests per second.

With a long enough key length, a brute force attack takes too long to be worth the attacker's efforts.

5.7.1.6 Guessing Attack

This type of attack is where an attacker attempts to simply “guess” the key. As an attack it is identical to the brute force attack (see Section 5.7.1.5 on page 647) where the odds of success depend on the length of the key.

5.7.1.7 Quantum Computer Attack

To break an n-bit key, a quantum computer [83] (NMR, Optical, or Caged Atom) containing n qubits embedded in an appropriate algorithm must be built. The quantum computer effectively exists in 2^(n) simultaneous coherent states. The trick is to extract the right coherent state without causing any decoherence. To date this has been achieved with a 2 qubit system (which exists in 4 coherent states). It is thought possible to extend this to 6 qubits (with 64 simultaneous coherent states) within a few years.

Unfortunately, every additional qubit halves the relative strength of the signal representing the key. This rapidly becomes a serious impediment to key retrieval, especially with the long keys. used in cryptographically secure systems.

As a result, attacks on a cryptographically secure key (e.g. 160 bits) using a Quantum Computer are likely not to be feasible and it is extremely unlikely that quantum computers will have achieved more than 50 or so qubits within the commercial lifetime of the authentication chips. Even using a 50 qubit quantum computer, 2¹¹⁰ tests are required to crack a 160 bit key.

5.7.1.8 Purposeful Error Attack

With certain algorithms, attackers can gather valuable information from the results of a bad input. This can range from the error message text to the time taken for the error to be generated.

A simple example is that of a userid/password scheme. If the error message usually says “Bad userid”, then when an attacker gets a message saying “Bad password” instead, then they know that the userid is correct. If the message always says “Bad userid/password” then much less information is given to the attacker. A more complex example is that of the recent published method of cracking encryption codes from secure web sites [41]. The attack involves sending particular messages to a server and observing the error message responses. The responses give enough information to learn the keys—even the lack of a response gives some information.

An example of algorithmic time can be seen with an algorithm that returns an error as soon as an erroneous bit is detected in the input message. Depending on hardware implementation, it may be a simple method for the attacker to time the response and alter each bit one by one depending on the time taken for the error response, and thus obtain the key. Certainly in a chip implementation the time taken can be observed with far greater accuracy than over the Internet.

5.7.1.9 Birthday Attack

This attack is named after the famous “birthday paradox” (which is not actually a paradox at all). The odds of one person sharing a birthday with another, is 1 in 365 (not counting leap years). Therefore there must be 183 people in a room for the odds to be more than 50% that one of them shares your birthday. However, there only needs to be 23 people in a room for there to be more than a 50% chance that any two share a birthday, as shown in the following relation:

${Prob} = {{1 - \frac{nPr}{n^{r}}} = {{1 - \frac{365{P23}}{365^{23}}} \approx 0.507}}$

Birthday attacks are common attacks against hashing algorithms, especially those algorithms that combine hashing with digital signatures.

If a message has been generated and already signed, an attacker must search for a collision message that hashes to the same value (analogous to finding one person who shares your birthday). However, if the attacker can generate the message, the birthday attack comes into play. The attacker searches for two messages that share the same hash value (analogous to any two people sharing a birthday), only one message is acceptable to the person signing it, and the other is beneficial for the attacker. Once the person has signed the original message the attacker simply claims now that the person signed the alternative message—mathematically there is no way to tell which message was the original, since they both hash to the same value.

Assuming a brute force attack is the only way to determine a match, the weakening of an n-bit key by the birthday attack is 2^(n/2). A key length of 128 bits that is susceptible to the birthday attack has an effective length of only 64 bits.

5.7.1.10 Chaining Attack

These are attacks made against the chaining nature of hash functions. They focus on the compression function of a hash function. The idea is based on the fact that a hash function generally takes arbitrary length input and produces a constant length output by processing the input n bits at a time. The output from one block is used as the chaining variable set into the next block. Rather than finding a collision against an entire input, the idea is that given an input chaining variable set, to find a substitute block that will result in the same output chaining variables as the proper message.

The number of choices for a particular block is based on the length of the block. If the chaining variable is c bits, the hashing function behaves like a random mapping, and the block length is b bits, the number of such b-bit blocks is approximately 2^(b)/2^(c). The challenge for finding a substitution block is that such blocks are a sparse subset of all possible blocks.

For SHA-1, the number of 512 bit blocks is approximately 2⁵¹²/2¹⁶⁰, or 2³⁵². The chance of finding a block by brute force search is about 1 in 2¹⁶⁰.

5.7.1.11 Substitution with a Complete Lookup Table

If the number of potential messages sent to the chip is small, then there is no need for a clone manufacturer to crack the key. Instead, the clone manufacturer could incorporate a ROM in their chip that had a record of all of the responses from a genuine chip to the codes sent by the system. The larger the key, and the larger the response, the more space is required for such a lookup table.

5.7.1.12 Substitution with a Sparse Lookup Table

If the messages sent to the chip are somehow predictable, rather than effectively random, then the clone manufacturer need not provide a complete lookup table. For example:

-   -   If the message is simply a serial number, the clone manufacturer         need simply provide a lookup table that contains values for past         and predicted future serial numbers. There are unlikely to be         more than 10⁹ of these.     -   If the test code is simply the date, then the clone manufacturer         can produce a lookup table using the date as the address.     -   If the test code is a pseudo-random number using either the         serial number or the date as a seed, then the clone manufacturer         just needs to crack the pseudo-random number generator in the         system. This is probably not difficult, as they have access to         the object code of the system. The clone manufacturer would then         produce a content addressable memory (or other sparse array         lookup) using these codes to access stored authentication codes.         5.7.1.13 Differential Cryptanalysis

Differential cryptanalysis describes an attack where pairs of input streams are generated with known differences, and the differences in the encoded streams are analyzed.

Existing differential attacks are heavily dependent on the structure of S boxes, as used in DES and other similar algorithms. Although other algorithms such as HMAC-SHA1 have no S boxes, an attacker can undertake a differential-like attack by undertaking statistical analysis of:

-   -   Minimal-difference inputs, and their corresponding outputs     -   Minimal-difference outputs, and their corresponding inputs

Most algorithms were strengthened against differential cryptanalysis once the process was described. This is covered in the specific sections devoted to each cryptographic algorithm. However some recent algorithms developed in secret have been broken because the developers had not considered certain styles of differential attacks [94] and did not subject their algorithms to public scrutiny.

5.7.1.14 Message Substitution Attacks

In certain protocols, a man-in-the-middle can substitute part or all of a message. This is where a real authentication chip is plugged into a reusable clone chip within the consumable. The clone chip intercepts all messages between the system and the authentication chip, and can perform a number of substitution attacks.

Consider a message containing a header followed by content. An attacker may not be able to generate a valid header, but may be able to substitute their own content, especially if the valid response is something along the lines of “Yes, I received your message”. Even if the return message is “Yes, I received the following message . . . ”, the attacker may be able to substitute the original message before sending the acknowledgment back to the original sender.

Message Authentication Codes were developed to combat message substitution attacks.

5.7.1.15 Reverse Engineering the Key Generator

If a pseudo-random number generator is used to generate keys, there is the potential for a clone manufacture to obtain the generator program or to deduce the random seed used. This was the way in which the security layer of the Netscape browser program was initially broken [33].

5.7.1.16 Bypassing the Authentication Process

It may be that there are problems in the authentication protocols that can allow a bypass of the authentication process altogether. With these kinds of attacks the key is completely irrelevant, and the attacker has no need to recover it or deduce it.

Consider an example of a system that authenticates at power-up, but does not authenticate at any other time. A reusable consumable with a clone authentication chip may make use of a real authentication chip. The clone authentication chip uses the real chip for the authentication call, and then simulates the real authentication chip's state data after that.

Another example of bypassing authentication is if the system authenticates only after the consumable has been used. A clone authentication chip can accomplish a simple authentication bypass by simulating a loss of connection after the use of the consumable but before the authentication protocol has completed (or even started).

One infamous attack known as the “Kentucky Fried Chip” hack [2] involved replacing a microcontroller chip for a satellite TV system. When a subscriber stopped paying the subscription fee, the system would send out a “disable” message. However the new micro-controller would simply detect this message and not pass it on to the consumer's satellite TV system.

5.7.1.17 Garrote/Bribe Attack

If people know the key, there is the possibility that they could tell someone else. The telling may be due to coercion (bribe, garrote etc.), revenge (e.g. a disgruntled employee), or simply for principle. These attacks are usually cheaper and easier than other efforts at deducing the key. As an example, a number of people claiming to be involved with the development of the (now defunct) Divx standard for DVD claimed (before the standard was rejected by consumers) that they would like to help develop Divx specific cracking devices—out of principle.

5.7.2 Physical Attacks

The following attacks assume implementation of an authentication mechanism in a silicon chip that the attacker has physical access to. The first attack, Reading ROM, describes an attack when keys are stored in ROM, while the remaining attacks assume that a secret key is stored in Flash memory.

5.7.2.1 Reading ROM

If a key is stored in ROM it can be read directly. A ROM can thus be safely used to hold a public key (for use in asymmetric cryptography), but not to hold a private key. In symmetric cryptography, a ROM is completely insecure. Using a copyright text (such as a haiku) as the key is not sufficient, because we are assuming that the cloning of the chip is occurring in a country where intellectual property is not respected.

5.7.2.2 Reverse Engineering of Chip

Reverse engineering of the chip is where an attacker opens the chip and analyzes the circuitry. Once the circuitry has been analyzed the inner workings of the chip's algorithm can be recovered. Lucent Technologies have developed an active method [4] known as TOBIC (Two photon OBIC, where OBIC stands for Optical Beam Induced Current), to image circuits. Developed primarily for static RAM analysis, the process involves removing any back materials, polishing the back surface to a mirror finish, and then focusing light on the surface. The excitation wavelength is specifically chosen not to induce a current in the IC.

A Kerckhoffs in the nineteenth century made a fundamental assumption about cryptanalysis: if the algorithm's inner workings are the sole secret of the scheme, the scheme is as good as broken [39]. He stipulated that the secrecy must reside entirely in the key. As a result, the best way to protect against reverse engineering of the chip is to make the inner workings irrelevant.

5.7.2.3 Usurping the Authentication Process

It must be assumed that any clone manufacturer has access to both the system and consumable designs.

If the same channel is used for communication between the system and a trusted system authentication chip, and a non-trusted consumable authentication chip, it may be possible for the non-trusted chip to interrogate a trusted authentication chip in order to obtain the “correct answer”. If this is so, a clone manufacturer would not have to determine the key. They would only have to trick the system into using the responses from the system authentication chip.

The alternative method of usurping the authentication process follows the same method as the logical attack described in Section 5.7.1.16 on page 652, involving simulated loss of contact with the system whenever authentication processes take place, simulating power-down etc.

5.7.2.4 Modification of System

This kind of attack is where the system itself is modified to accept clone consumables. The attack may be a change of system ROM, a rewiring of the consumable, or, taken to the extreme case, a completely clone system.

Note that this kind of attack requires each individual system to be modified, and would most likely require the owner's consent. There would usually have to be a clear advantage for the consumer to undertake such a modification, since it would typically void warranty and would most likely be costly. An example of such a modification with a clear advantage to the consumer is a software patch to change fixed-region DVD players into region-free DVD players (although it should be noted that this is not to use clone consumables, but rather originals from the same companies simply targeted for sale in other countries).

5.7.2.5 Direct Viewing of Chip Operation by Conventional Probing

If chip operation could be directly viewed using an STM (Scanning Tunnelling Microscope) or an electron beam, the keys could be recorded as they are read from the internal non-volatile memory and loaded into work registers.

These forms of conventional probing require direct access to the top or front sides of the IC while it is powered.

5.7.2.6 Direct Viewing of the Non-Volatile Memory

If the chip were sliced so that the floating gates of the Flash memory were exposed, without discharging them, then the key could probably be viewed directly using an STM or SKM (Scanning Kelvin Microscope).

However, slicing the chip to this level without discharging the gates is probably impossible. Using wet etching, plasma etching, ion milling (focused ion beam etching), or chemical mechanical polishing will almost certainly discharge the small charges present on the floating gates.

5.7.2.7 Viewing the Light Bursts Caused by State Changes

Whenever a gate changes state, a small amount of infrared energy is emitted. Since silicon is transparent to infrared, these changes can be observed by looking at the circuitry from the underside of a chip. While the emission process is weak, it is bright enough to be detected by highly sensitive equipment developed for use in astronomy. The technique [92], developed by IBM, is called PICA (Picosecond Imaging Circuit Analyzer). If the state of a register is known at time t, then watching that register change over time will reveal the exact value at time t+n, and if the data is part of the key, then that part is compromised.

5.7.2.8 Viewing the Keys Using an SEPM

A non-invasive testing device, known as a Scanning Electric Potential Microscope (SEPM), allows the direct viewing of charges within a chip [37]. The SEPM has a tungsten probe that is placed a few micrometers above the chip, with the probe and circuit forming a capacitor. Any AC signal flowing beneath the probe causes displacement current to flow through this capacitor. Since the value of the current change depends on the amplitude and phase of the AC signal, the signal can be imaged. If the signal is part of the key, then that part is compromised.

5.7.2.9 Monitoring EMI

Whenever electronic circuitry operates, faint electromagnetic signals are given off. Relatively inexpensive equipment can monitor these signals and could give enough information to allow an attacker to deduce the keys.

5.7.2.10 Viewing I_(dd) Fluctuations

Even if keys cannot be viewed, there is a fluctuation in current whenever registers change state. If there is a high enough signal to noise ratio, an attacker can monitor the difference in I_(dd) that may occur when programming over either a high or a low bit. The change in I_(dd) can reveal information about the key. Attacks such as these have already been used to break smart cards [46].

5.7.2.11 Differential Fault Analysis

This attack assumes introduction of a bit error by ionization, microwave radiation, or environmental stress. In most cases such an error is more likely to adversely affect the chip (e.g. cause the program code to crash) rather than cause beneficial changes which would reveal the key. Targeted faults such as ROM overwrite, gate destruction etc. are far more likely to produce useful results.

5.7.2.12 Clock Glitch Attacks

Chips are typically designed to properly operate within a certain clock speed range. Some attackers attempt to introduce faults in logic by running the chip at extremely high clock speeds or introduce a clock glitch at a particular time for a particular duration [1]. The idea is to create race conditions where the circuitry does not function properly. An example could be an AND gate that (because of race conditions) gates through Input₁ all the time instead of the AND of Input, and Input₂.

If an attacker knows the internal structure of the chip, they can attempt to introduce race conditions at the correct moment in the algorithm execution, thereby revealing information about the key (or in the worst case, the key itself).

5.7.2.13 Power Supply Attacks

Instead of creating a glitch in the clock signal, attackers can also produce glitches in the power supply where the power is increased or decreased to be outside the working operating voltage range. The net effect is the same as a clock glitch—introduction of error in the execution of a particular instruction. The idea is to stop the CPU from XORing the key, or from shifting the data one bit-position etc. Specific instructions are targeted so that information about the key is revealed.

5.7.2.14 Overwriting ROM

Single bits in a ROM can be overwritten using a laser cutter microscope [1], to either 1 or 0 depending on the sense of the logic. If the ROM contains instructions, it may be a simple matter for an attacker to change a conditional jump to a non-conditional jump, or perhaps change the destination of a register transfer. If the target instruction is chosen carefully, it may result in the key being revealed.

5.7.2.15 Modifying EEPROM/Flash

These attacks fall into two categories:

-   -   those similar to the ROM attacks except that the laser cutter         microscope technique can be used to both set and reset         individual bits. This gives much greater scope in terms of         modification of algorithms.     -   Electron beam programming of floating gates. As described in         [89] and [32], a focused electron beam can change a gate by         depositing electrons onto it. Damage to the rest of the circuit         can be avoided, as described in [31].         5.7.2.16 Gate Destruction

Anderson and Kuhn described the rump session of the 1997 workshop on Fast Software Encryption (1], where Biham and Shamir presented an attack on DES. The attack was to use a laser cutter to destroy an individual gate in the hardware implementation of a known block cipher (DES). The net effect of the attack was to force a particular bit of a register to be “stuck”. Biham and Shamir described the effect of forcing a particular register to be affected in this way—the least significant bit of the output from the round function is set to 0. Comparing the 6 least significant bits of the left half and the right half can recover several bits of the key. Damaging a number of chips in this way can reveal enough information about the key to make complete key recovery easy.

An encryption chip modified in this way will have the property that encryption and decryption will no longer be inverses.

5.7.2.17 Overwrite Attacks

Instead of trying to read the Flash memory, an attacker may simply set a single bit by use of a laser cutter microscope. Although the attacker doesn't know the previous value, they know the new value. If the chip still works, the bit's original state must be the same as the new state. If the chip doesn't work any longer, the bit's original state must be the logical NOT of the current state. An attacker can perform this attack on each bit of the key and obtain the n-bit key using at most n chips (if the new bit matched the old bit, a new chip is not required for determining the next bit).

5.7.2.18 Test Circuitry Attack

Most chips contain test circuitry specifically designed to check for manufacturing defects. This includes BIST (Built In Self Test) and scan paths. Quite often the scan paths and test circuitry includes access and readout mechanisms for all the embedded latches. In some cases the test circuitry could potentially be used to give information about the contents of particular registers.

Test circuitry is often disabled once the chip has passed all manufacturing tests, in some cases by blowing a specific connection within the chip. A determined attacker, however, can reconnect the test circuitry and hence enable it.

5.7.2.19 Memory Remnants

Values remain in RAM long after the power has been removed [35], although they do not remain long enough to be considered non-volatile. An attacker can remove power once sensitive information has been moved into RAM (for example working registers), and then attempt to read the value from RAM. This attack is most useful against security systems that have regular RAM chips. A classic example is cited by [1], where a security system was designed with an automatic power-shut-off that is triggered when the computer case is opened. The attacker was able to simply open the case, remove the RAM chips, and retrieve the key because the values persisted.

5.7.2.20 Chip Theft Attack

If there are a number of stages in the lifetime of an authentication chip, each of these stages must be examined in terms of ramifications for security should chips be stolen. For example, if information is programmed into the chip in stages, theft of a chip between stages may allow an attacker to have access to key information or reduced efforts for attack. Similarly, if a chip is stolen directly after manufacture but before programming, does it give an attacker any logical or physical advantage?

5.7.2.21 Trojan Horse Attack

At some stage the authentication chips must be programmed with a secret key. Suppose an attacker builds a clone authentication chip and adds it to the pile of chips to be programmed. The attacker has especially built the clone chip so that it looks and behaves just like a real authentication chip, but will give the key out to the attacker when a special attacker-known command is issued to the chip. Of course the attacker must have access to the chip after the programming has taken place, as well as physical access to add the Trojan horse authentication chip to the genuine chips.

6 Requirements

Existing solutions to the problem of authenticating consumables have typically relied on patents covering physical packaging. However this does not stop home refill operations or clone manufacture in countries with weak industrial property protection. Consequently a much higher level of protection is required.

The authentication mechanism is therefore built into an authentication chip that is embedded in the consumable and allows a system to authenticate that consumable securely and easily. Limiting ourselves to the system authenticating consumables (we don't consider the consumable authenticating the system), two levels of protection can be considered:

Presence Only Authentication:

-   -   This is where only the presence of an authentication chip is         tested. The authentication chip can be removed and used in other         consumables as long as be used indefinitely.         Consumable Lifetime Authentication:     -   This is where not only is the presence of the authentication         chip tested for, but also the authentication chip must only last         the lifetime of the consumable. For the chip to be re-used it         must be completely erased and reprogrammed.

The two levels of protection address different requirements. We are primarily concerned with Consumable Lifetime authentication in order to prevent cloned versions of high volume consumables. In this case, each chip should hold secure state information about the consumable being authenticated. It should be noted that a Consumable Lifetime authentication chip could be used in any situation requiring a Presence Only authentication chip.

Requirements for authentication, data storage integrity and manufacture are considered separately. The following sections summarize requirements of each.

6.1 Authentication

The authentication requirements for both Presence Only and Consumable Lifetime authentication are restricted to the case of a system authenticating a consumable. We do not consider bi-directional authentication where the consumable also authenticates the system. For example, it is not necessary for a valid toner cartridge to ensure it is being used in a valid photocopier.

For Presence Only authentication, we must be assured that an authentication chip is physically present. For Consumable Lifetime authentication we also need to be assured that state data actually came from the authentication chip, and that it has not been altered en route. These issues cannot be separated—data that has been altered has a new source, and if the source cannot be determined, the question of alteration cannot be settled.

It is not enough to provide an authentication method that is secret, relying on a home-brew security method that has not been scrutinized by security experts. The primary requirement therefore is to provide authentication by means that have withstood the scrutiny of experts.

The authentication scheme used by the authentication chip should be resistant to defeat by logical means. Logical types of attack are extensive, and attempt to do one of three things:

-   -   Bypass the authentication process altogether     -   Obtain the secret key by force or deduction, so that any         question can be answered     -   Find enough about the nature of the authenticating questions and         answers in order to, without the key, give the right answer to         each question.

The logical attack styles and the forms they take are detailed in Section 5.7.1 on page 646.

The algorithm should have a flat keyspace, allowing any random bit string of the required length to be a possible key. There should be no weak keys.

6.2 Data Storage Integrity

Although authentication protocols take care of ensuring data integrity in communicated messages, data storage integrity is also required. Two kinds of data must be stored within the authentication chip:

-   -   Authentication data, such as secret keys     -   Consumable state data, such as serial numbers, and media         remaining etc.

The access requirements of these two data types differ greatly. The authentication chip therefore requires a storage/access control mechanism that allows for the integrity requirements of each type.

6.2.1 Authentication Data

Authentication data must remain confidential. It needs to be stored in the chip during a manufacturing/programming stage of the chip's life, but from then on must not be permitted to leave the chip. It must be resistant to being read from non-volatile memory. The authentication scheme is responsible for ensuring the key cannot be obtained by deduction, and the manufacturing process is responsible for ensuring that the key cannot be obtained by physical means.

The size of the authentication data memory area must be large enough to hold the necessary keys and secret information as mandated by the authentication protocols.

6.2.2 Consumable State Data

Consumable state data can be divided into the following types. Depending on the application, there will be different numbers of each of these types of data items.

-   -   Read Only     -   ReadWrite     -   Decrement Only

-   Read Only data needs to be stored in the chip during a     manufacturing/programming stage of the chip's life, but from then on     should not be allowed to change. Examples of Read Only data items     are consumable batch numbers and serial numbers.

-   ReadWrite data is changeable state information, for example, the     last time the particular consumable was used. ReadWrite data items     can be read and written an unlimited number of times during the     lifetime of the consumable. They can be used to store any state     information about the consumable. The only requirement for this data     is that it needs to be kept in non-volatile memory. Since an     attacker can obtain access to a system (which can write to ReadWrite     data), any attacker can potentially change data fields of this type.     This data type should not be used for secret information, and must     be considered insecure.

-   Decrement Only data is used to count down the availability of     consumable resources. A photocopier's toner cartridge, for example,     may store the amount of toner remaining as a Decrement Only data     item. An ink cartridge for a color printer may store the amount of     each ink color as a Decrement Only data item, requiring 3 (one for     each of Cyan, Magenta, and Yellow), or even as many as 5 or 6     Decrement Only data items. The requirement for this kind of data     item is that once programmed with an initial value at the     manufacturing/programming stage, it can only reduce in value. Once     it reaches the minimum value, it cannot decrement any further. The     Decrement Only data item is only required by Consumable Lifetime     authentication.

Note that the size of the consumable state data storage required is only for that information required to be authenticated. Information which would be of no use to an attacker, such as ink color-curve characteristics or ink viscosity do not have to be stored in the secure state data memory area of the authentication chip.

6.3 Manufacture

The authentication chip must have a low manufacturing cost in order to be included as the authentication mechanism for low cost consumables.

The authentication chip should use a standard manufacturing process, such as Flash. This is necessary to:

-   -   Allow a great range of manufacturing location options     -   Use well-defined and well-behaved technology     -   Reduce cost

Regardless of the authentication scheme used, the circuitry of the authentication part of the chip must be resistant to physical attack. Physical attack comes in four main ways, although the form of the attack can vary:

-   -   Bypassing the authentication chip altogether     -   Physical examination of chip while in operation (destructive and         non-destructive)     -   Physical decomposition of chip     -   Physical alteration of chip

The physical attack styles and the forms they take are detailed in Section 5.7.2 on page 652. Ideally, the chip should be exportable from the USA, so it should not be possible to use an authentication chip as a secure encryption device. This is low priority requirement since there are many companies in other countries able to manufacture the authentication chips. In any case, the export restrictions from the USA may change.

Authentication

7 Introduction

Existing solutions to the problem of authenticating consumables have typically relied on physical patents on packaging. However this does not stop home refill operations or clone manufacture in countries with weak industrial property protection. Consequently a much higher level of protection is required.

It is not enough to provide an authentication method that is secret, relying on a home-brew security method that has not been scrutinized by security experts. Security systems such as Netscape's original proprietary system and the GSM Fraud Prevention Network used by cellular phones are examples where design secrecy caused the vulnerability of the security [33][33]. Both security systems were broken by conventional means that would have been detected if the companies had followed an open design process. The solution is to provide authentication by means that have withstood the scrutiny of experts.

In this section, we examine a number of protocols that can be used for consumables authentication. We only use security methods that are publicly described, using known behaviors in this new way. Readers should be familiar with the concepts and terms described in Section 5 on page 629. We avoid the Zero Knowledge Proof protocol since it is patented.

For all protocols, the security of the scheme relies on a secret key, not a secret algorithm. In the nineteenth century, A Kerckhoffs made a fundamental assumption about cryptanalysis: if the algorithm's inner workings are the sole secret of the scheme, the scheme is as good as broken [39]. He stipulated that the secrecy must reside entirely in the key. As a result, the best way to protect against reverse engineering of any authentication chip is to make the algorithmic inner workings irrelevant (the algorithm of the inner workings must still be must be valID, but not the actual secret).

The QA Chip is a programmable device, and can therefore be setup with an application-specific program together with an application-specific set of protocols. This section describes the following sets of protocols:

-   -   single key single memory vector     -   multiple key single memory vector     -   multiple key multiple memory vector

These protocols refer to the number of valid keys that an QA Chip knows about, and the size of data required to be stored in the chip.

From these protocols it is straightforward to construct protocol sets for the single key multiple memory vector case (of course the multiple memory vector can be considered to be. and multiple key single memory vector. Other protocol sets can also be defined as necessary. Of course multiple memory vector can be conveniently

All the protocols rely on a time-variant challenge (i.e. the challenge is different each time), where the response depends on the challenge and the secret. The challenge involves a random number so that any observer will not be able to gather useful information about a subsequent identification.

8 Single Key Single Memory Vector

8.1 Protocol Background

This protocol set is provided for two reasons:

-   -   the other protocol sets defined in this document are simply         extensions of this one; and     -   it is useful in its own right

The single key protocol set is useful for applications where only a single key is required. Note that there can be many consumables and systems, but there is only a single key that connects them all. Examples include:

-   -   car and keys. A car and the car-key share a single key. There         can be multiple sets of car-keys, each effectively cut to the         same key. A company could have a set of cars, each with the same         key. Any of the car-keys could then be used to drive any of the         cars.     -   printer and ink cartridge. All printers of a certain model use         the same ink cartridge, with printer and cartridge sharing only         a single key. Note that to introduce a new printer model that         accepts the old ink cartridge the new model would need the same         key as the old model. See the multiple-key protocols for         alternative solutions to this problem.         8.2 Requirements of Protocol

Each QA Chip contains the following values:

-   -   K The secret key for calculating F_(K)[X]. K must not be stored         directly in the QA Chip.         -   Instead, each chip needs to store a random number R_(K)             (different for each chip), K⊕R_(K), and             K⊕R_(K). The stored K⊕R_(K) can be XORed with R_(K) to             obtain the real K.         -   Although             KβR_(K) must be stored to protect against differential             attacks, it is not used.     -   R Current random number used to ensure time varying messages.         Each chip instance must be seeded with a different initial         value. Changes for each signature generation.     -   M Memory vector of QA Chip.     -   P 2 element array of access permissions for each part of M.         Entry 0 holds access permissions for non-authenticated writes to         M (no key required). Entry 1 holds access permissions for         authenticated writes to M (key required). Permission choices for         each part of M are Read Only, Read/Write, and Decrement Only.     -   C 3 constants used for generating signatures. C₁, C₂, and C₃ are         constants that pad out a submessage to a hashing boundary, and         all 3 must be different.         -   Each QA Chip contains the following private function:     -   S_(K)[X] Internal function only. Returns S_(K)[X], the result of         applying a digital signature function S to X based upon key K.         The digital signature must be long enough to counter the chances         of someone generating a random signature. The length depends on         the signature scheme chosen, although the scheme chosen for the         QA Chip is HMAC-SHA1 (see Section 13 on page 691), and therefore         the length of the signature is 160 bits.

Additional functions are required in certain QA Chips, but these are described as required.

8.3 Reads of M

In this case, we have a trusted chip (ChipT) connected to a System. The System wants to authenticate an object that contains a non-trusted chip (ChipA). In effect, the System wants to know that it can securely read a memory vector (M) from ChipA: to be sure that ChipA is valid and that M has not been altered.

The protocol requires the following publicly available function in ChipA:

-   -   Read[X] Advances R, and returns R, M, S_(K)[X|R|C₁|M]. The time         taken to calculate the signature must not be based on the         contents of X, R, M, or K.

The protocol requires the following publicly available functions in ChipT:

-   -   Random         Returns R (does not advance R).     -   Test[)(, Y, Z] Advances R and returns 1 if S_(K)[R|X|C₁|Y]=Z.         Otherwise returns 0. The time taken to calculate and compare         signatures must be independent of data content.

To authenticate ChipA and read ChipA's memory M:

-   a. System calls ChipT's Random function; -   b. ChipT returns R_(T) to System; -   c. System calls ChipA's Read function, passing in the result from b; -   d. ChipA updates R_(A), then calculates and returns R_(A), M_(A),     S_(K)[R_(T)|R_(A)|C₁|M_(A)]; -   e. System calls ChipT's Test function, passing in R_(A), M_(A),     S_(K)[R_(T)|R_(A)|C₁|M_(A)]; -   f. System checks response from ChipT. If the response is 1, then     ChipA is considered authentic. If 0, ChipA is considered invalid.

The data flow for read authentication is shown in FIG. 334.

The protocol allows System to simply pass data from one chip to another, with no special processing. The protection relies on ChipT being trusted, even though System does not know K.

When ChipT is physically separate from System (eg is chip on a board connected to System) System must also occassionally (based on system clock for example) call ChipT's Test function with bad data, expecting a 0 response. This is to prevent someone from inserting a fake ChipT into the system that always returns 1 for the Test function.

8.4 Writes

In this case, the System wants to update M in some chip referred to as ChipU. This can be non-authenticated (for example, anyone is allowed to count down the amount of consumable remaining), or authenticated (for example, replenishing the amount of consumable remaining).

8.4.1 Non-Authenticated Writes

This is the most frequent type of write, and takes place between the System/consumable during normal everyday operation. In this kind of write, System wants to change M in a way that doesn't require special authorization. For example, the System could be decrementing the amount of consumable remaining. Although System does not need to know K or even have access to a trusted chip, System must follow a non-authenticated write by an authenticated read if it needs to know that the write was successful.

The protocol requires the following publicly available function:

-   -   Write[X] Writes X over those parts of M subject to P₀ and the         existing value for M.

To authenticate a write of M_(new) to ChipA's memory M:

-   a. System calls ChipU's Write function, passing in M_(new); -   b. The authentication procedure for a Read is carried out (see     Section 8.3 on page 664); -   c. If ChipU is authentic and M_(new)=M returned in b, the write     succeeded. If not, it failed.     8.4.2 Authenticated Writes

In this kind of write, System wants to change Chip U's M in an authorized way, without being subject to the permissions that apply during normal operation (P₀). For example, the consumable may be at a refilling station and the normally Decrement Only section of M should be updated to include the new valid consumable. In this case, the chip whose M is being updated must authenticate the writes being generated by the external System and in addition, apply permissions P₁ to ensure that only the correct parts of M are updated.

In this transaction protocol, the System's chip is referred to as ChipS, and the chip being updated is referred to as ChipU. Each chip distrusts the other.

The protocol requires the following publicly available functions in ChipU:

-   -   Read[X] Advances R, and returns R, M, S_(K)[X|R|C₁|M]. The time         taken to calculate the signature must be identical for all         inputs.     -   WriteA[X, Y, Z]Returns 1, advances R, and replaces M by Y         subject to P₁ only if S_(K)[R|X|C₁|Y]=Z. Otherwise returns 0.         The time taken to calculate and compare signatures must be         independent of data content. This function is identical to         ChipT's Test function except that it additionally writes Y over         those parts of M subject to P₁ when the signature matches.

Authenticated writes require that the System has access to a ChipS that is capable of generating appropriate signatures. ChipS requires the following variables and function:

-   -   CountRemaining Part of M that contains the number of signatures         that ChipS is allowed to generate. Decrements with each         successful call to SignM and SignP.         -   Permissions in ChipS's P₀ for this part of M needs to be             ReadOnly once ChipS has been setup. Therefore CountRemaining             can only be updated by another ChipS that will perform             updates to that part of M (assuming ChipS's P₁ allows that             part of M to be updated).     -   Q Part of M that contains the write permissions for updating         ChipU's M. By adding Q to ChipS we allow different ChipSs that         can update different parts of M_(U). Permissions in ChipS's P₀         for this part of M needs to be ReadOnly once ChipS has been         setup. Therefore Q can only be updated by another ChipS that         will perform updates to that part of M.     -   SignM[V,W,X,Y,Z] Advances R, decrements CountRemaining and         returns R, Z_(QX) (Z applied to X with permissions Q), followed         by S_(K)[W|R|C₁|Z_(QX)] only if S_(K)[V|W|C₁|X]=Y and         CountRemaining >0. Otherwise returns all 0s. The time taken to         calculate and compare signatures must be independent of data         content.

To update ChipU's M vector:

-   a. System calls ChipU's Read function, passing in 0 as the input     parameter; -   b. ChipU produces R_(U), M_(U), S_(K)[0|R_(U)|C₁|M_(U)] and returns     these to System; -   c. System calls ChipS's SignM function, passing in 0 (as used in a),     R_(U), M_(U), S_(K)[0|R_(U)|C₁|M_(U)], and M_(D) (the desired vector     to be written to ChipU); -   d. ChipS produces R_(S), M_(QD) (processed by running MD against     M_(U) using Q) and S_(K)[R_(U)|R_(S)|C₁|M_(QD)] if the inputs were     valID, and 0 for all outputs if the inputs were not valid. -   e. If values returned in d are non zero, then ChipU is considered     authentic. System can then call ChipU's WriteA function with these     values. -   f. ChipU should return a 1 to indicate success. A 0 should only be     returned if the data generated by ChipS is incorrect (e.g. a     transmission error).

The data flow for authenticated writes is shown in FIG. 335.

Note that Q in ChipS is part of ChipS's M. This allows a user to set up ChipS with a permission set for upgrades. This should be done to ChipS and that part of M designated by P₀ set to ReadOnly before ChipS is programmed with K_(U). If K_(S) is programmed with K_(U) first, there is a risk of someone obtaining a half-setup ChipS and changing all of M_(U) instead of only the sections specified by Q.

The same is true of CountRemaining. The CountRemaining value needs to be setup (including making it ReadOnly in P₀) before ChipS is programmed with K_(U). ChipS is therefore programmed to only perform a limited number of SignM operations (thereby limiting compromise exposure if a ChipS is stolen). Thus ChipS would itself need to be upgraded with a new CountRemaining every so often.

8.4.3 Updating Permissions for Future Writes

In order to reduce exposure to accidental and malicious attacks on P and certain parts of M, only authorized users are allowed to update P. Writes to P are the same as authorized writes to M, except that they update P_(n) instead of M. Initially (at manufacture), P is set to be Read/Write for all parts of M. As different processes fill up different parts of M, they can be sealed against future change by updating the permissions. Updating a chip's P₀ changes permissions for unauthorized writes, and updating P₁ changes permissions for authorized writes.

P_(n) is only allowed to change to be a more restrictive form of itself. For example, initially all parts of M have permissions of Read/Write. A permission of Read/Write can be updated to Decrement Only or Read Only. A permission of Decrement Only can be updated to become Read Only. A Read Only permission cannot be further restricted.

In this transaction protocol, the System's chip is referred to as ChipS, and the chip being updated is referred to as ChipU. Each chip distrusts the other.

The protocol requires the following publicly available functions in ChipU:

-   -   Random         Returns R (does not advance R).     -   SetPermission[n,X,Y,Z] Advances R, and updates P_(n) according         to Y and returns 1 followed by the resultant P_(n) only if         S_(K)[R|X|Y|C₂]=Z. Otherwise returns 0. P_(n) can only become         more restricted. Passing in 0 for any permission leaves it         unchanged (passing in Y=0 returns the current P_(n)).

Authenticated writes of permissions require that the System has access to a ChipS that is capable of generating appropriate signatures. ChipS requires the following variables and function:

-   -   CountRemaining Part of M that contains the number of signatures         that ChipS is allowed to generate. Decrements with each         successful call to SignM and SignP.         -   Permissions in ChipS's P₀ for this part of M needs to be             ReadOnly once ChipS has been setup. Therefore CountRemaining             can only be updated by another ChipS that will perform             updates to that part of M (assuming ChipS's P₁ allows that             part of M to be updated).     -   SignP[X,Y] Advances R, decrements CountRemaining and returns R         and S_(K)[X|R|Y|C₂] only if CountRemaining>0. Otherwise returns         all 0s. The time taken to calculate and compare signatures must         be independent of data content.

To update ChipU's P_(n):

-   a. System calls ChipU's Random function; -   b. ChipU returns R_(U) to System; -   c. System calls ChipS's SignP function, passing in R_(U) and P_(D)     (the desired P to be written to ChipU); -   d. ChipS produces R_(S) and S_(K)[R_(U)|R_(S)|P_(D)|C₂] if it is     still permitted to produce signatures. -   e. If values returned in d are non zero, then System can then call     ChipU's SetPermission function with the desired n, R_(S), P_(D) and     S_(K)[R_(U)|R_(S)|P_(D)|C₂]. -   f. ChipU verifies the received signature against     S_(K)[R_(U)|R_(S)|P_(D)|C₂] and applies P_(D) to P_(n) if the     signature matches -   g. System checks 1st output parameter. 1=success, 0=failure.

The data flow for authenticated writes to permissions is shown in FIG. 336 below.

8.5 Programming K

In this case, we have a factory chip (ChipF) connected to a System. The System wants to program the key in another chip (ChipP). System wants to avoid passing the new key to ChipP in the clear, and also wants to avoid the possibility of the key-upgrade message being replayed on another ChipP (even if the user doesn't know the key).

The protocol assumes that ChipF and ChipP already share a secret key K_(old). This key is used to ensure that only a chip that knows K_(old) can set K_(new).

The protocol requires the following publicly available functions in ChipP:

-   -   Random         Returns R (does not advance R).     -   ReplaceKey[X, Y, Z] Replaces K by S_(Kold)[R|X|C₃]⊕Y, advances         R, and returns 1 only if S_(Kold)[X|Y|C₃]=Z. Otherwise         returns 0. The time taken to calculate signatures and compare         values must be identical for all inputs.

And the following data and function in ChipF:

-   -   CountRemaining Part of M with contains the number of signatures         that ChipF is allowed to generate. Decrements with each         successful call to GetProgramKey. Permissions in P for this part         of M needs to be ReadOnly once ChipF has been setup. Therefore         can only be updated by a ChipS that has authority to perform         updates to that part of M.     -   K_(new) The new key to be transferred from ChipF to ChipP. Must         not be visible.     -   SetPartialKey[X,Y] If word X of K_(new) has not yet been set,         set word X of K_(new) to Y and return 1. Otherwise return 0.         This function allows K_(new) to be programmed in multiple steps,         thereby allowing different people or systems to know different         parts of the key (but not the whole K_(new)). K_(new) is stored         in ChipF's flash memory. Since there is a small number of         ChipFs, it is theoretically not necessary to store the inverse         of K_(new), but it is stronger protection to do so.     -   GetProgramKey[X] Advances R_(F), decrements CountRemaining,         outputs R_(F), the encrypted key S_(Kold)[X|R_(F)|C₃]⊕K_(new)         and a signature of the first two outputs plus C₃ if         CountRemaining>0. Otherwise outputs 0. The time to calculate the         encrypted key & signature must be identical for all inputs.

To update P's key:

-   a. System calls ChipP's Random function; -   b. ChipP returns R_(P) to System; -   c. System calls ChipF's GetProgramKey function, passing in the     result from b; -   d. ChipF updates R_(F), then calculates and returns R_(F),     S_(Kold)[R_(P)|R_(F)|C₃]⊕K_(new), and     S_(Kold)[R_(F)|S_(Kold)[R_(P)|R_(F)|C₃]⊕K_(new)|C₃]; -   e. If the response from d is not 0, System calls ChipP's ReplaceKey     function, passing in the response from d; -   f. System checks response from ChipP. If the response is 1, then     K_(P) has been correctly updated to K_(new). If the response is 0,     K_(P) has not been updated.

The data flow for key updates is shown in FIG. 337.

Note that K_(new) is never passed in the open. An attacker could send its own R_(P), but cannot produce S_(Kold)[R_(P)|R_(F)|C₃] without K_(old). The third parameter, a signature, is sent to ensure that ChipP can determine if either of the first two parameters have been changed en route.

CountRemaining needs to be setup in M_(F) (including making it ReadOnly in P) before ChipF is programmed with K_(P). ChipF should therefore be programmed to only perform a limited number of GetProgramKey operations (thereby limiting compromise exposure if a ChipF is stolen). An authorized ChipS can be used to update this counter if neccesary (see Section 8.4 on page 665).

8.5.1 Chicken and Egg

Of course, for the Program Key protocol to work, both ChipF and ChipP must both know K_(old). Obviously both chips had to be programmed with K_(old), and thus K_(old) can be thought of as an older K_(new):K_(old) can be placed in chips if another ChipF knows K_(older), and so on.

Although this process allows a chain of reprogramming of keys, with each stage secure, at some stage the very first key (K_(first)) must be placed in the chips. K_(first) is in fact programmed with the chip's microcode at the manufacturing test station as the last step in manufacturing test. K_(first) can be a manufacturing batch key, changed for each batch or for each customer etc, and can have as short a life as desired. Compromising K_(first) need not result in a complete compromise of the chain of K_(S).

9 Multiple Key Single Memory Vector

9.1 Protocol Background

This protocol set is an extension to the single key single memory vector protocol set, and is provided for two reasons:

-   -   the multiple key multiple memory vector protocol set defined in         this document is simply extensions of this one; and     -   it is useful in its own right

The multiple key protocol set is typically useful for applications where there are multiple types of systems and consumables, and they need to work with each other in various ways. This is typically in the following situations:

-   -   when different systems want to share some consumables, but not         others. For example printer models may share some ink cartridges         and not share others.     -   when there are different owners of data in M. Part of the memory         vector may be owned by one company (eg the speed of the printer)         and another may be owned by another (eg the serial number of the         chip). In this case a given key K_(n) needs to be able to write         to a given part of M, and other keys K_(n) need to be disallowed         from writing to these same areas.         9.2 Requirements of Protocol

Each QA Chip contains the following values:

-   N The maximum number of keys known to the chip. -   K_(N) Array of N secret keys used for calculating F_(Kn)[X] where     K_(n) is the nth element of the array. Each K_(n) must not be stored     directly in the QA Chip. Instead, each chip needs to store a single     random number R_(K) (different for each chip), K_(n)⊕R_(K), and     K_(n)⊕R_(K). The stored K_(n)⊕R_(K) can be XORed with R_(K) to     obtain the real K_(n). Although     K_(n)⊕R_(K) must be stored to protect against differential attacks,     it is not used. -   R Current random number used to ensure time varying messages. Each     chip instance must be seeded with a different initial value. Changes     for each signature generation. -   M Memory vector of QA Chip. A fixed part of M contains N in ReadOnly     form so users of the chip can know the number of keys known by the     chip. -   P N+1 element array of access permissions for each part of M. Entry     0 holds access permissions for non-authenticated writes to M (no key     required). Entries 1 to N+1 hold access permissions for     authenticated writes to M, one for each K. Permission choices for     each part of M are Read Only, Read/Write, and Decrement Only. -   C 3 constants used for generating signatures. C₁, C₂, and C₃ are     constants that pad out a submessage to a hashing boundary, and all 3     must be different.

Each QA Chip contains the following private function:

-   S_(Kn)[N,X] Internal function only. Returns S_(Kn)[X], the result of     applying a digital signature function S to X based upon the     appropriate key K_(n). The digital signature must be long enough to     counter the chances of someone generating a random signature. The     length depends on the signature scheme chosen, although the scheme     chosen for the QA Chip is HMAC-SHA1 (see Section 13 on page 691),     and therefore the length of the signature is 160 bits.

Additional functions are required in certain QA Chips, but these are described as required.

9.3 Reads

As with the single key scenario, we have a trusted chip (ChipT) connected to a System. The System wants to authenticate an object that contains a non-trusted chip (ChipA). In effect, the System wants to know that it can securely read a memory vector (M) from ChipA: to be sure that ChipA is valid and that M has not been altered.

The protocol requires the following publicly available functions:

-   -   Random         Returns R (does not advance R).     -   Read[n, X] Advances R, and returns R, M, S_(Kn)[X|R|C₁|M]. The         time taken to calculate the signature must not be based on the         contents of X, R, M, or K.     -   Test[n,X, Y, Z] Advances R and returns 1 if S_(Kn)[R|X|C₁|Y]=Z.         Otherwise returns 0. The time taken to calculate and compare         signatures must be independent of data content.

To authenticate ChipA and read ChipA's memory M:

-   a. System calls ChipT's Random function; -   b. ChipT returns R_(T) to System; -   c. System calls ChipA's Read function, passing in some key number n1     and the result from b; -   d. ChipA updates R_(A), then calculates and returns R_(A), M_(A),     S_(KAn1)[R_(T)|R_(A)|C₁|M_(A)]; -   e. System calls ChipT's Test function, passing in n2, R_(A), M_(A),     S_(KAn1)[R_(T)|R_(A)|C₁|M_(A)]; -   f. System checks response from ChipT. If the response is 1, then     ChipA is considered authentic. If 0, ChipA is considered invalid.

The choice of n1 and n2 must be such that ChipA's K_(n1)=ChipT's K_(n2).

The data flow for read authentication is shown in FIG. 338.

The protocol allows System to simply pass data from one chip to another, with no special processing. The protection relies on ChipT being trusted, even though System does not know K.

When ChipT is physically separate from System (eg is chip on a board connected to System) System must also occassionally (based on system clock for example) call ChipT's Test function with bad data, expecting a 0 response. This is to prevent someone from inserting a fake ChipT into the system that always returns 1 for the Test function.

It is important that n1 is chosen by System. Otherwise ChipA would need to return N_(A) sets of signatures for each read, since ChipA does not know which of the keys will satisfy ChipT. Similarly, system must also choose n2, so it can potentially restrict the number of keys in ChipT that are matched against (otherwise ChipT would have to match against all its keys). This is important in order to restrict how different keys are used. For example, say that ChipT contains 6 keys, keys 0-2 are for various printer-related upgrades, and keys 3-6 are for inks. ChipA contains say 4 keys, one key for each printer model. At power-up, System goes through each of chipA's keys 0-3, trying each out against ChipT's keys 3-6. System doesn't try to match against ChipT's keys 0-2. Otherwise knowledge of a speed-upgrade key could be used to provide ink QA Chip chips. This matching needs to be done only once (eg at power up). Once matching keys are found, System can continue to use those key numbers.

Since System needs to know N_(T) and N_(A), part of M is used to hold N (eg in Read Only form), and the system can obtain it by calling the Read function, passing in key 0.

9.4 Writes

As with the single key scenario, the System wants to update M in ChipU. As before, this can be done in a non-authenticated and authenticated way.

9.4.1 Non-Authenticated Writes

This is the most frequent type of write, and takes place between the System/consumable during normal everyday operation. In this kind of write, System wants to change M subject to P. For example, the System could be decrementing the amount of consumable remaining. Although System does not need to know any of the K_(S) or even have access to a trusted chip to perform the write, System must follow a non-authenticated write by an authenticated read if it needs to know that the write was successful.

The protocol requires the following publicly available function:

-   Write[X] Writes X over those parts of M subject to P₀ and the     existing value for M.

To authenticate a write of M_(new) to ChipA's memory M:

-   a. System calls ChipU's Write function, passing in M_(new); -   b. The authentication procedure for a Read is carried out (see     Section 9.3 on page 671); -   c. If ChipU is authentic and M_(new)=M returned in b, the write     succeeded. If not, it failed.     9.4.2 Authenticated Writes

In this kind of write, System wants to change Chip U's M in an authorized way, without being subject to the permissions that apply during normal operation (P₀). For example, the consumable may be at a refilling station and the normally Decrement Only section of M should be updated to include the new valid consumable. In this case, the chip whose M is being updated must authenticate the writes being generated by the external System and in addition, apply the appropriate permission for the key to ensure that only the correct parts of M are updated. Having a different permission for each key is required as when multiple keys are involved, all keys should not necessarily be given open access to M. For example, suppose M contains printer speed and a counter of money available for franking. A ChipS that updates printer speed should not be capable of updating the amount of money. Since P₀ is used for non-authenticated writes, each K_(n) has a corresponding permission P_(n+1) that determines what can be updated in an authenticated write.

In this transaction protocol, the System's chip is referred to as ChipS, and the chip being updated is referred to as ChipU. Each chip distrusts the other.

The protocol requires the following publicly available functions in ChipU:

-   -   Read[n, X] Advances R, and returns R, M, S_(Kn)[X|R|C₁|M]. The         time taken to calculate the signature must not be based on the         contents of X, R, M, or K.     -   WriteA[n, X, Y, Z] Advances R, replaces M by Y subject to         P_(n+1), and returns 1 only if S_(Kn)[R|X|C₁|Y]=Z. Otherwise         returns 0. The time taken to calculate and compare signatures         must be independent of data content. This function is identical         to ChipT's Test function except that it additionally writes Y         subject to P_(n+1) to its M when the signature matches.

Authenticated writes require that the System has access to a ChipS that is capable of generating appropriate signatures. ChipS requires the following variables and function:

-   CountRemaining Part of M that contains the number of signatures that     ChipS is allowed to generate. Decrements with each successful call     to SignM and SignP.     -   -   Permissions in ChipS's P₀ for this part of M needs to be             ReadOnly once ChipS has been setup. Therefore CountRemaining             can only be updated by another ChipS that will perform             updates to that part of M (assuming ChipS's P allows that             part of M to be updated). -   Q Part of M that contains the write permissions for updating     ChipU's M. By adding Q to ChipS we allow different ChipSs that can     update different parts of M_(U). Permissions in ChipS's P₀ for this     part of M needs to be ReadOnly once ChipS has been setup. Therefore     Q can only be updated by another ChipS that will perform updates to     that part of M. -   SignM[n,V,W,X,Y,Z] Advances R, decrements CountRemaining and returns     R, Z_(QX) (Z applied to X with permissions Q), S_(Kn)[W|R|C₁|Z_(QX)]     only if Y=S_(Kn)[V|W|C₁|X] and CountRemaining>0. Otherwise returns     all 0s. The time taken to calculate and compare signatures must be     independent of data content.

To update ChipU's M vector:

-   a. System calls ChipU's Read function, passing in n1 and 0 as the     input parameters; -   b. ChipU produces R_(U), M_(U), S_(Kn1)[0|R_(U)|C₁|M_(U)] and     returns these to System; -   c. System calls ChipS's SignM function, passing in n2 (the key to be     used in ChipS), 0 (as used in a), R_(U), M_(U),     S_(Kn1)[0|R_(U)|C₁|M_(U)], and M_(D) (the desired vector to be     written to ChipU); -   d. ChipS produces R_(S), M_(QD) (processed by running M_(D) against     M_(U) using Q) and S_(Kn2)[R_(U)|R_(S)|C₁|M_(QD)] if the inputs were     valID, and 0 for all outputs if the inputs were not valid. -   e. If values returned in d are non zero, then ChipU is considered     authentic. System can then call ChipU's WriteA function with these     values from d. -   f. ChipU should return a 1 to indicate success. A 0 should only be     returned if the data generated by ChipS is incorrect (e.g. a     transmission error).

The choice of n1 and n2 must be such that ChipU's K_(n1)=ChipS's K_(n2).

The data flow for authenticated writes is shown in FIG. 339 below.

Note that Q in ChipS is part of ChipS's M. This allows a user to set up ChipS with a permission set for upgrades. This should be done to ChipS and that part of M designated by P₀ set to ReadOnly before ChipS is programmed with K_(U). If K_(S) is programmed with K_(U) first, there is a risk of someone obtaining a half-setup ChipS and changing all of M_(U) instead of only the sections specified by Q.

In addition, CountRemaining in ChipS needs to be setup (including making it ReadOnly in P_(S)) before ChipS is programmed with K_(U). ChipS should therefore be programmed to only perform a limited number of SignM operations (thereby limiting compromise exposure if a ChipS is stolen). Thus ChipS would itself need to be upgraded with a new CountRemaining every so often.

9.4.3 Updating Permissions for Future Writes

In order to reduce exposure to accidental and malicious attacks on P (and certain parts of M), only authorized users are allowed to update P. Writes to P are the same as authorized writes to M, except that they update P_(n) instead of M. Initially (at manufacture), P is set to be Read/Write for all parts of M. As different processes fill up different parts of M, they can be sealed against future change by updating the permissions. Updating a chip's P₀ changes permissions for unauthorized writes, and updating P_(n+1) changes permissions for authorized writes with key K_(n).

P_(n) is only allowed to change to be a more restrictive form of itself. For example, initially all parts of M have permissions of Read/Write. A permission of Read/Write can be updated to Decrement Only or Read Only. A permission of Decrement Only can be updated to become Read Only. A Read Only permission cannot be further restricted.

In this transaction protocol, the System's chip is referred to as ChipS, and the chip being updated is referred to as ChipU. Each chip distrusts the other.

The protocol requires the following publicly available functions in ChipU:

-   Random     Returns R (does not advance R). -   SetPermission[n,p,X,Y,Z] Advances R, and updates P_(p) according to     Y and returns 1 followed by the resultant P_(p) only if     S_(Kn)[R|X|Y|C₂]=Z. Otherwise returns 0. P_(p) can only become more     restricted. Passing in 0 for any permission leaves it unchanged     (passing in Y=0 returns the current P_(p)).

Authenticated writes of permissions require that the System has access to a ChipS that is capable of generating appropriate signatures. ChipS requires th following variables and function:

-   CountRemaining Part of M that contains the number of signatures that     ChipS is allowed to generate. Decrements with each successful call     to SignM and SignP.     -   -   Permissions in ChipS's P₀ for this part of M needs to be             ReadOnly once ChipS has been setup. Therefore CountRemaining             can only be updated by another ChipS that will perform             updates to that part of M (assuming ChipS's P_(n) allows             that part of M to be updated). -   SignP[n,X,Y] Advances R, decrements CountRemaining and returns R and     S_(Kn)[X|R|Y|C₂] only if CountRemaining>0. Otherwise returns all 0s.     The time taken to calculate and compare signatures must be     independent of data content.

To update ChipU's P_(n):

-   a. System calls ChipU's Random function; -   b. ChipU returns R_(U) to System; -   c. System calls ChipS's SignP function, passing in n1, R_(U) and     P_(D) (the desired P to be written to ChipU); -   d. ChipS produces R_(S) and S_(Kn1)[R_(U)|R_(S)|P_(D)|C₂] if it is     still permitted to produce signatures. -   e. If values returned in d are non zero, then System can then call     ChipU's SetPermission function with n2, the desired permission entry     p, R_(S), P_(D) and S_(Kn1)[R_(U)|R_(S)|P_(D)|C₂]. -   f. ChipU verifies the received signature against     S_(Kn2)[R_(U)|R_(S)|P_(D)|C₂] and applies P_(D) to P_(n) if the     signature matches -   g. System checks 1st output parameter. 1=success, 0=failure.

The choice of n1 and n2 must be such that ChipU's K_(n1)=ChipS's K_(n2).

The data flow for authenticated writes to permissions is shown in FIG. 340 below.

9.4.4 Protecting M in a Multiple Key System

To protect the appropriate part of M, the SetPermission function must be called after the part of M has been set to the desired value.

For example, if adding a serial number to an area of M that is currently ReadWrite so that noone is permitted to update the number again:

-   -   the Write function is called to write the serial number to M     -   SetPermission is called for n={1, . . . , N} to set that part of         M to be ReadOnly for authorized writes using key n−1.     -   SetPermission is called for 0 to set that part of M to be         ReadOnly for non-authorized writes

For example, adding a consumable value to M such that only keys 1-2 can update it, and keys 0, and 3-N cannot:

-   -   the Write function is called to write the amount of consumable         to M     -   SetPermission is called for n={1, 4, 5, . . . , N−1} to set that         part of M to be ReadOnly for authorized writes using key n−1.         This leaves keys 1 and 2 with ReadWrite permissions.     -   SetPermission is called for 0 to set that part of M to be         DecrementOnly for non-authorized writes. This allows the amount         of consumable to decrement.

It is possible for someone who knows a key to further restrict other keys, but it is not in anyone's interest to do so.

9.5 Programming K

In this case, we have a factory chip (ChipF) connected to a System. The System wants to program the key in another chip (ChipP). System wants to avoid passing the new key to ChipP in the clear, and also wants to avoid the possibility of the key-upgrade message being replayed on another ChipP (even if the user doesn't know the key).

The protocol is a simple extension of the single key protocol in that it assumes that ChipF and ChipP already share a secret key K_(old). This key is used to ensure that only a chip that knows K_(old) can set K_(new).

The protocol requires the following publicly available functions in ChipP:

-   -   Random         Returns R (does not advance R).     -   ReplaceKey[n, X, Y, Z] Replaces K_(n) by S_(Kn)[R|X|C₃]⊕Y,         advances R, and returns 1 only if S_(Kn)[X|Y|C₃]=Z. Otherwise         returns 0. The time taken to calculate signatures and compare         values must be identical for all inputs.

And the following data and functions in ChipF:

-   -   CountRemaining Part of M with contains the number of signatures         that ChipF is allowed to generate. Decrements with each         successful call to GetProgramKey.         -   Permissions in P for this part of M needs to be ReadOnly             once ChipF has been setup. Therefore can only be updated by             a ChipS that has authority to perform updates to that part             of M.     -   K_(new) The new key to be transferred from ChipF to ChipP. Must         not be visible.     -   SetPartialKey[X,Y] If word X of K_(new) has not yet been set,         set word X of K_(new) to Y and return 1. Otherwise return 0.         This function allows K_(new) to be programmed in multiple steps,         thereby allowing different people or systems to know different         parts of the key (but not the whole K_(new)). K_(new) is stored         in ChipF's flash memory. Since there is a small number of         ChipFs, it is theoretically not necessary to store the inverse         of K_(new), but it is stronger protection to do so.     -   GetProgramKey[n, X] Advances R_(F), decrements CountRemaining,         outputs R_(F), the encrypted key S_(Kn)[X|R_(F)|C₃]⊕K_(new) and         a signature of the first two outputs plus C₃ if         CountRemaining>0. Otherwise outputs 0. The time to calculate the         encrypted key & signature must be identical for all inputs.

To update P's key:

-   a. System calls ChipP's Random function; -   b. ChipP returns R_(P) to System; -   c. System calls ChipF's GetProgramKey function, passing in n1 (the     desired key to use) and the result from b; -   d. ChipF updates R_(F), then calculates and returns R_(F),     S_(Kn1)[R_(P)|R_(F)|C₃]⊕K_(new), and     S_(Kn1)[R_(F)|S_(Kn1)[R_(P)|R_(F)|C₃]⊕K_(new)|C₃]; -   e. If the response from d is not 0, System calls ChipP's ReplaceKey     function, passing in n2 (the key to use in ChipP) and the response     from d; -   f. System checks response from ChipP. If the response is 1, then     K_(Pn2) has been correctly updated to K_(new). If the response is 0,     K_(Pn2) has not been updated.

The choice of n1 and n2 must be such that ChipF's K_(n1)=ChipP's K_(n2).

The data flow for key updates is shown in FIG. 341 below.

Note that K_(new) is never passed in the open. An attacker could send its own R_(P), but cannot produce S_(Kn1)[R_(P)|R_(F)|C₃] without K_(n1). The signature based on K_(new) is sent to ensure that ChipP will be able to determine if either of the first two parameters have been changed en route.

CountRemaining needs to be setup in M_(F) (including making it ReadOnly in P) before ChipF is programmed with K_(P). ChipF should therefore be programmed to only perform a limited number of GetProgramKey operations (thereby limiting compromise exposure if a ChipF is stolen). An authorized ChipS can be used to update this counter if neccesary (see Section 9.4 on page 673).

9.5.1 Chicken and Egg

As with the single key protocol, for the Program Key protocol to work, both ChipF and ChipP must both know K_(old). Obviously both chips had to be programmed with K_(old), and thus K_(old) can be thought of as an older K_(new):K_(old) can be placed in chips if another ChipF knows K_(older), and so on.

Although this process allows a chain of reprogramming of keys, with each stage secure, at some stage the very first key (K_(first)) must be placed in the chips. K_(first) is in fact programmed with the chip's microcode at the manufacturing test station as the last step in manufacturing test. K_(first) can be a manufacturing batch key, changed for each batch or for each customer etc, and can have as short a life as desired. Compromising K_(first) need not result in a complete compromise of the chain of Ks.

Depending on the reprogramming requirements, K_(first) can be the same or different for all K_(n).

10 Multiple Keys Multiple Memory Vectors

10.1 Protocol Background

This protocol set is a slight restriction of the multiple key single memory vector protocol set, and is the expected protocol. It is a restriction in that M has been optimized for Flash memory utilization.

M is broken into multiple memory vectors (semi-fixed and variable components) for the purposes of optimizing flash memory utilization. Typically M contains some parts that are fixed at some stage of the manufacturing process (eg a batch number, serial number etc), and once set, are not ever updated. This information does not contain the amount of consumable remaining, and therefore is not read or written to with any great frequency.

We therefore define M₀ to be the M that contains the frequently updated sections, and the remaining Ms to be rarely written to. Authenticated writes only write to M₀, and non-authenticated writes can be directed to a specific M_(n). This reduces the size of permissions that are stored in the QA Chip (since key-based writes are not required for Ms other than M₀). It also means that M₀ and the remaining Ms can be manipulated in different ways, thereby increasing flash memory longevity.

10.2 Requirements of Protocol

Each QA Chip contains the following values:

-   N The maximum number of keys known to the chip. -   T The number of vectors M is broken into. -   K_(N) Array of N secret keys used for calculating F_(Kn)[X] where     K_(n) is the nth element of the array.

Each K_(n) must not be stored directly in the QA Chip. Instead, each chip needs to store a single random number R_(K) (different for each chip), K_(n)⊕R_(K), and

K_(n)⊕R_(K). The stored K_(n)⊕R_(K) can be XORed with R_(K) to obtain the real K_(n). Although

K_(n)⊕R_(K) must be stored to protect against differential attacks, it is not used.

-   R Current random number used to ensure time varying messages. Each     chip instance must be seeded with a different initial value. Changes     for each signature generation. -   M_(T) Array of T memory vectors. Only M₀ can be written to with an     authorized write, while all Ms can be written to in an unauthorized     write. Writes to M₀ are optimized for Flash usage, while updates to     any other M_(n) are expensive with regards to Flash utilization, and     are expected to be only performed once per section of M_(n). M₁     contains T and N in ReadOnly form so users of the chip can know     these two values. -   P_(T+N) T+N element array of access permissions for each part of M.     Entries n={0 . . . T−1} hold access permissions for     non-authenticated writes to M_(n) (no key required). Entries n={T to     T+N−1}hold access permissions for authenticated writes to M₀ for     K_(n). Permission choices for each part of M are Read Only,     Read/Write, and Decrement Only. -   C 3 constants used for generating signatures. C₁, C₂, and C₃ are     constants that pad out a submessage to a hashing boundary, and all 3     must be different.

Each QA Chip contains the following private function:

-   S_(Kn)[N,X] Internal function only. Returns S_(Kn)[X], the result of     applying a digital signature function S to X based upon the     appropriate key K_(n). The digital signature must be long enough to     counter the chances of someone generating a random signature. The     length depends on the signature scheme chosen, although the scheme     chosen for the QA Chip is HMAC-SHA1, and therefore the length of the     signature is 160 bits.

Additional functions are required in certain QA Chips, but these are described as required.

10.3 Reads

As with the previous scenarios, we have a trusted chip (ChipT) connected to a System. The System wants to authenticate an object that contains a non-trusted chip (ChipA). In effect, the System wants to know that it can securely read a memory vector (M_(t)) from ChipA: to be sure that ChipA is valid and that M has not been altered.

The protocol requires the following publicly available functions:

-   -   Random         Returns R (does not advance R).     -   Read[n, t, X] Advances R, and returns R, M_(t),         S_(Kn)[X|R|C₁|M_(t)]. The time taken to calculate the signature         must not be based on the contents of X, R, M_(t), or K. If t is         invalid, the function assumes t=0.     -   Test[n,X, Y, Z] Advances R and returns 1 if S_(Kn)[R|X|C₁|Y]=Z.         Otherwise returns 0. The time taken to calculate and compare         signatures must be independent of data content.

To authenticate ChipA and read ChipA's memory M:

-   a. System calls ChipT's Random function; -   b. ChipT returns R_(T) to System; -   c. System calls ChipA's Read function, passing in some key number     n1, the desired M number t, and the result from b; -   d. ChipA updates R_(A), then calculates and returns R_(A), M_(At),     S_(KAn1)[R_(T)|R_(A)|C₁|M_(At)]; -   e. System calls ChipT's Test function, passing in n2, R_(A), M_(At),     S_(KAn1)[R_(T)|R_(A)|C₁|M_(At)]; -   f. System checks response from ChipT. If the response is 1, then     ChipA is considered authentic. If 0, ChipA is considered invalid.

The choice of n1 and n2 must be such that ChipA's K_(n1)=ChipT's K_(n2).

The data flow for read authentication is shown in FIG. 342 below.

The protocol allows System to simply pass data from one chip to another, with no special processing. The protection relies on ChipT being trusted, even though System does not know K.

When ChipT is physically separate from System (eg is chip on a board connected to System) System must also occassionally (based on system clock for example) call ChipT's Test function with bad data, expecting a 0 response. This is to prevent someone from inserting a fake ChipT into the system that always returns 1 for the Test function.

It is important that n1 is chosen by System. Otherwise ChipA would need to return NA sets of signatures for each read, since ChipA does not know which of the keys will satisfy ChipT. Similarly, system must also choose n2, so it can potentially restrict the number of keys in ChipT that are matched against (otherwise ChipT would have to match against all its keys). This is important in order to restrict how different keys are used. For example, say that ChipT contains 6 keys, keys 0-2 are for various printer-related upgrades, and keys 3-6 are for inks. ChipA contains say 4 keys, one key for each printer model. At power-up, System goes through each of chipA's keys 0-3, trying each out against ChipT's keys 3-6. System doesn't try to match against ChipT's keys 0-2. Otherwise knowledge of a speed-upgrade key could be used to provide ink QA Chip chips. This matching needs to be done only once (eg at power up). Once matching keys are found, System can continue to use those key numbers.

Since System needs to know N_(T), N_(A), and T_(A), part of M₁ is used to hold N (eg in Read Only form), and the system can obtain it by calling the Read function, passing in key 0 and t=1.

10.4 Writes

As with the previous scenarios, the System wants to update M_(t) in ChipU. As before, this can be done in a non-authenticated and authenticated way.

10.4.1 Non-Authenticated Writes

This is the most frequent type of write, and takes place between the System/consumable during normal everyday operation for M₀, and during the manufacturing process for M_(t).

In this kind of write, System wants to change M subject to P. For example, the System could be decrementing the amount of consumable remaining. Although System does not need to know and of the Ks or even have access to a trusted chip to perform the write, System must follow a non-authenticated write by an authenticated read if it needs to know that the write was successful.

The protocol requires the following publicly available function:

-   -   Write[t, X] Writes X over those parts of M_(t) subject to P_(t)         and the existing value for M.

To authenticate a write of M_(new) to ChipA's memory M:

-   a. System calls ChipU's Write function, passing in M_(new); -   b. The authentication procedure for a Read is carried out (see     Section 9.3 on page 671); -   c. If ChipU is authentic and M_(new)=M returned in b, the write     succeeded. If not, it failed.     10.4.2 Authenticated Writes

In the multiple memory vectors protocol, only M₀ can be written to an authenticated way. This is because only M₀ is considered to have components that need to be upgraded.

In this kind of write, System wants to change Chip U's M₀ in an authorized way, without being subject to the permissions that apply during normal operation. For example, the consumable may be at a refilling-station and the normally Decrement Only section of M₀ should be updated to include the new valid consumable. In this case, the chip whose M₀ is being updated must authenticate the writes being generated by the external System and in addition, apply the appropriate permission for the key to ensure that only the correct parts of M₀ are updated. Having a different permission for each key is required as when multiple keys are involved, all keys should not necessarily be given open access to M₀. For example, suppose M₀ contains printer speed and a counter of money available for franking. A ChipS that updates printer speed should not be capable of updating the amount of money. Since P_(0 . . . T−1) is used for non-authenticated writes, each K_(n) has a corresponding permission P_(T+n) that determines what can be updated in an authenticated write.

In this transaction protocol, the System's chip is referred to as ChipS, and the chip being updated is referred to as ChipU. Each chip distrusts the other.

The protocol requires the following publicly available functions in ChipU:

-   -   Read[n, t, X] Advances R, and returns R, M_(t),         S_(Kn)[X|R|C₁|M_(t)]. The time taken to calculate the signature         must not be based on the contents of X, R, M_(t), or K.     -   WriteA[n, X, Y, Z] Advances R, replaces M₀ by Y subject to         P_(T+n), and returns 1 only if S_(Kn)[R|X|C₁|Y]=Z. Otherwise         returns 0. The time taken to calculate and compare signatures         must be independent of data content. This function is identical         to ChipT's Test function except that it additionally writes Y         subject to P_(T+n) to its M when the signature matches.

Authenticated writes require that the System has access to a ChipS that is capable of generating appropriate signatures. ChipS requires the following variables and function:

-   -   CountRemaining Part of M that contains the number of signatures         that ChipS is allowed to generate. Decrements with each         successful call to SignM and SignP. Permissions in ChipS's         P_(0..T−1) for this part of M needs to be ReadOnly once ChipS         has been setup. Therefore CountRemaining can only be updated by         another ChipS that will perform updates to that part of M         (assuming ChipS's P allows that part of M to be updated).     -   Q Part of M that contains the write permissions for updating         ChipU's M. By adding Q to ChipS we allow different ChipSs that         can update different parts of M_(U). Permissions in ChipS's         P_(0..T−1) for this part of M needs to be ReadOnly once ChipS         has been setup. Therefore Q can only be updated by another ChipS         that will perform updates to that part of M.     -   SignM[n,V,W,X,Y,Z] Advances R, decrements CountRemaining and         returns R, Z_(QX) (Z applied to X with permissions Q),         S_(Kn)[W|R|C₁|Z_(QX) only if Y=S_(Kn)[V|W|C₁|X] and         CountRemaining>0. Otherwise returns all 0s. The time taken to         calculate and compare signatures must be independent of data         content.

To update ChipU's M vector:

-   a. System calls ChipU's Read function, passing in n1, 0 and 0 as the     input parameters; -   b. ChipU produces R_(U), M_(U0), S_(Kn1)[0|R_(U)|C₁|M_(U0)] and     returns these to System; -   c. System calls ChipS's SignM function, passing in n2 (the key to be     used in ChipS), 0 (as used in a), R_(U), M_(U0),     S_(Kn1)[0|R_(U)|C₁|M_(U0)], and M_(D) (the desired vector to be     written to ChipU); -   d. ChipS produces R_(S), M_(QD) (processed by running M_(D) against     M_(U0) using Q) and S_(Kn2)[R_(U)|R_(S)|C₁|M_(QD)] if the inputs     were valID, and 0 for all outputs if the inputs were not valid. -   e. If values returned in d are non zero, then ChipU is considered     authentic. System can then call ChipU's WriteA function with these     values from d. -   f. ChipU should return a 1 to indicate success. A 0 should only be     returned if the data generated by ChipS is incorrect (e.g. a     transmission error).

The choice of n1 and n2 must be such that ChipU's K_(n1)=ChipS's K_(n2).

The data flow for authenticated writes is shown in FIG. 343 below.

Note that Q in ChipS is part of ChipS's M. This allows a user to set up ChipS with a permission set for upgrades. This should be done to ChipS and that part of M designated by P_(0..T−1) set to ReadOnly before ChipS is programmed with K_(U). If K_(S) is programmed with K_(U) first, there is a risk of someone obtaining a half-setup ChipS and changing all of M_(U) instead of only the sections specified by Q.

In addition, CountRemaining in ChipS needs to be setup (including making it ReadOnly in P_(S)) before ChipS is programmed with K_(U). ChipS should therefore be programmed to only perform a limited number of SignM operations (thereby limiting compromise exposure if a ChipS is stolen). Thus ChipS would itself need to be upgraded with a new CountRemaining every so often.

10.4.3 Updating Permissions for Future Writes

In order to reduce exposure to accidental and malicious attacks on P (and certain parts of M), only authorized users are allowed to update P. Writes to P are the same as authorized writes to M, except that they update P_(n) instead of M. Initially (at manufacture), P is set to be Read/Write for all M. As different processes fill up different parts of M, they can be sealed against future change by updating the permissions. Updating a chip's P_(0..T−1) changes permissions for unauthorized writes to M_(n), and updating P_(T..T+N−1) changes permissions for authorized writes with key K_(n).

P_(n) is only allowed to change to be a more restrictive form of itself. For example, initially all parts of M have permissions of Read/Write. A permission of Read/Write can be updated to Decrement Only or Read Only. A permission of Decrement Only can be updated to become Read Only. A Read Only permission cannot be further restricted.

In this transaction protocol, the System's chip is referred to as ChipS, and the chip being updated is referred to as ChipU. Each chip distrusts the other.

The protocol requires the following publicly available functions in ChipU:

-   -   Random         Returns R (does not advance R).     -   SetPermission[n,p,X,Y,Z] Advances R, and updates P_(p) according         to Y and returns 1 followed by the resultant P_(p) only if         S_(Kn)[R|X|Y|C₂]=Z. Otherwise returns 0. P_(p) can only become         more restricted. Passing in 0 for any permission leaves it         unchanged (passing in Y=0 returns the current P_(p)).

Authenticated writes of permissions require that the System has access to a ChipS that is capable of generating appropriate signatures. ChipS requires the following variables and function:

-   -   CountRemaining Part of ChipS's M₀ that contains the number of         signatures that ChipS is allowed to generate. Decrements with         each successful call to SignM and SignP. Permissions in ChipS's         P_(0..T−1) for this part of M₀ needs to be ReadOnly once ChipS         has been setup. Therefore CountRemaining can only be updated by         another ChipS that will perform updates to that part of M₀         (assuming ChipS's P_(n) allows that part of M₀ to be updated).     -   SignP[n,X,Y] Advances R, decrements CountRemaining and returns R         and S_(Kn)[X|R|Y|C₂] only if CountRemaining>0. Otherwise returns         all 0s. The time taken to calculate and compare signatures must         be independent of data content.

To update ChipU's P_(n):

-   a. System calls ChipU's Random function; -   b. ChipU returns R_(U) to System; -   c. System calls ChipS's SignP function, passing in n1, R_(U) and     P_(D) (the desired P to be written to ChipU); -   d. ChipS produces R_(S) and S_(Kn1)[R_(U)|R_(S)|P_(D)|C₂] if it is     still permitted to produce signatures. -   e. If values returned in d are non zero, then System can then call     ChipU's SetPermission function with n2, the desired permission entry     p, R_(S), P_(D) and S_(Kn1)[R_(U)|R_(S)|P_(D)|C₂]. -   f. ChipU verifies the received signature against     S_(Kn2)[R_(U)|R_(S)|P_(D)|C₂] and applies P_(D) to P_(n) if the     signature matches -   g. System checks 1 st output parameter. 1=success, 0=failure.

The choice of n1 and n2 must be such that ChipU's K_(n1)=ChipS's K_(n2).

The data flow for authenticated writes to permissions is shown in FIG. 344 below.

10.4.4 Protecting M in a Multiple Key Multiple M System

To protect the appropriate part of M_(n) against unauthorized writes, call SetPermissions[n] for n=0 to T−1. To protect the appropriate part of M₀ against authorized writes with key n, call SetPermissions[T+n] for n=0 to N−1.

Note that only M₀ can be written in an authenticated fashion.

Note that the SetPermission function must be called after the part of M has been set to the desired value.

For example, if adding a serial number to an area of M₁ that is currently ReadWrite so that noone is permitted to update the number again:

-   -   the Write function is called to write the serial number to M₁     -   SetPermission(1) is called for to set that part of M to be         ReadOnly for non-authorized writes.

If adding a consumable value to M₀ such that only keys 1-2 can update it, and keys 0, and 3-N cannot:

-   -   the Write function is called to write the amount of consumable         to M     -   SetPermission is called for 0 to set that part of M₀ to be         DecrementOnly for non-authorized writes. This allows the amount         of consumable to decrement.     -   SetPermission is called for n={T, T+3, T+4 . . . , T+N−1} to set         that part of M₀ to be ReadOnly for authorized writes using all         but keys 1 and 2. This leaves keys 1 and 2 with ReadWrite         permissions to M₀.

It is possible for someone who knows a key to further restrict other keys, but it is not in anyone's interest to do so.

10.5 Programming K

This section is identical to the multiple key single memory vector (Section 9.5 on page 677). It is repeated here with mention to M₀ instead of M for CountRemaining.

In this case, we have a factory chip (ChipF) connected to a System. The System wants to program the key in another chip (ChipP). System wants to avoid passing the new key to ChipP in the clear, and also wants to avoid the possibility of the key-upgrade message being replayed on another ChipP (even if the user doesn't know the key).

The protocol is a simple extension of the single key protocol in that it assumes that ChipF and ChipP already share a secret key K_(old). This key is used to ensure that only a chip that knows K_(old) can set K_(new).

The protocol requires the following publicly available functions in ChipP:

-   -   Random         Returns R (does not advance R).     -   ReplaceKey[n, X, Y, Z] Replaces K_(n) by S_(Kn)[R|X|C₃]⊕Y,         advances R, and returns 1 only if S_(Kn)[X|Y|C₃]=Z. Otherwise         returns 0. The time taken to calculate signatures and compare         values must be identical for all inputs.

And the following data and functions in ChipF:

-   -   CountRemaining Part of M₀ with contains the number of signatures         that ChipF is allowed to generate. Decrements with each         successful call to GetProgramKey. Permissions in P for this part         of M₀ needs to be ReadOnly once ChipF has been setup. Therefore         can only be updated by a ChipS that has authority to perform         updates to that part of M₀.     -   K_(new) The new key to be transferred from ChipF to ChipP. Must         not be visible.     -   SetPartialKey[X,Y] If word X of K_(new) has not yet been set,         set word X of K_(new) to Y and return 1. Otherwise return 0.         This function allows K_(new) to be programmed in multiple steps,         thereby allowing different people or systems to know different         parts of the key (but not the whole K_(new)). K_(new) is stored         in ChipF's flash memory. Since there is a small number of         ChipFs, it is theoretically not necessary to store the inverse         of K_(new), but it is stronger protection to do so.     -   GetProgramKey[n, X] Advances R_(F), decrements CountRemaining,         outputs R_(F), the encrypted key S_(Kn)[X|R_(F)|C₃]⊕K_(new) and         a signature of the first two outputs plus C₃ if         CountRemaining>0. Otherwise outputs 0. The time to calculate the         encrypted key & signature must be identical for all inputs.

To update P's key:

-   a. System calls ChipP's Random function; -   b. ChipP returns R_(P) to System; -   c. System calls ChipF's GetProgramKey function, passing in n1 (the     desired key to use) and the result from b; -   d. ChipF updates R_(F), then calculates and returns R_(F),     S_(Kn1)[R_(P)|R_(F)|C₃]⊕K_(new), and     S_(Kn1)[R_(F)|S_(Kn1)[R_(P)|R_(F)|C₃]⊕K_(new)|C₃]; -   e. If the response from d is not 0, System calls ChipP's ReplaceKey     function, passing in n2 (the key to use in ChipP) and the response     from d; -   f. System checks response from ChipP. If the response is 1, then     K_(Pn2) has been correctly updated to K_(new). If the response is 0,     K_(Pn2) has not been updated.

The choice of n1 and n2 must be such that ChipF's K_(n1)=ChipP's K_(n2).

The data flow for key updates is shown in FIG. 345 below.

Note that K_(new) is never passed in the open. An attacker could send its own R_(P), but cannot produce S_(Kn1)[R_(P)|R_(F)|C₃] without K_(n1). The signature based on K_(new) is sent to ensure that ChipP will be able to determine if either of the first two parameters have been changed en route.

CountRemaining needs to be setup in M_(F0) (including making it ReadOnly in P) before ChipF is programmed with K_(P). ChipF should therefore be programmed to only perform a limited number of GetProgramKey operations (thereby limiting compromise exposure if a ChipF is stolen). An authorized ChipS can be used to update this counter if neccesary (see Section 9.4 on page 673).

10.5.1 Chicken and Egg

As with the single key protocol, for the Program Key protocol to work, both ChipF and ChipP must both know K_(old). Obviously both chips had to be programmed with K_(old), and thus K_(old) can be thought of as an older K_(new):K_(old) can be placed in chips if another ChipF knows K_(older), and so on.

Although this process allows a chain of reprogramming of keys, with each stage secure, at some stage the very first key (K_(first)) must be placed in the chips. K_(first) is in fact programmed with the chip's microcode at the manufacturing test station as the last step in manufacturing test. K_(first) can be a manufacturing batch key, changed for each batch or for each customer etc, and can have as short a life as desired. Compromising K_(first) need not result in a complete compromise of the chain of Ks. Depending on reprogramming requirements, K_(first) can be the same or different for all K_(n).

10.5.2 Security Note

Different ChipFs should have different R_(F) values to prevent K_(new) from being determined as follows: The attacker needs 2 ChipFs, both with the same R_(F) and K_(n) but different values for K_(new). By knowing K_(new1) the attacker can determine K_(new2). The size of R_(F) is 2¹⁶⁰, and assuming a lifespan of approximately 2³² R_(S), an attacker needs about 2⁶⁰ ChipFs with the same K_(n) to locate the correct chip. Given that there are likely to be only hundreds of ChipFs with the same K_(n), this is not a likely attack. The attack can be eliminated completely by making C₃ different per chip and transmitting it with the new signature.

11 Summary of Functions for All Protocols

All protocol sets, whether single key, multiple key, single M or multiple M, all rely on the same set of functions. The function set is listed here:

11.1 All Chips

Since every chip must act as ChipP, ChipA and potentially ChipU, all chips require the following functions:

-   -   Random     -   ReplaceKey     -   Read     -   Write     -   WriteA     -   SetPermissions         11.2 ChipT

Chips that are to be used as ChipT also require:

-   -   Test         11.3 ChipS

Chips that are to be used as ChipS also require either or both of:

-   -   SignM     -   SignP         11.4 ChipF

Chips that are to be used as ChipF also require:

-   -   SetPartialKey     -   GetProgramKey         12 Remote Upgrades         12.1 Basic Remote Upgrades

Regardless of the number of keys and the number of memory vectors, the use of authenticated reads and writes, and of replacing a new key without revealing K_(new) or K_(old) allows the possibility of remote upgrades of ChipU and ChipP. The upgrade typically involves a remote server and follows two basic steps:

-   a. During the first stage of the upgrade, the remote system     authenticates the user's system to ensure the user's system has the     setup that it claims to have. -   b. During the second stage of the upgrade, the user's system     authenticates the remote system to ensure that the upgrade is from a     trusted source.     12.1.1 User Requests Upgrade

The user requests that he wants to upgrade. This can be done by running a specific upgrade application on the user's computer, or by visiting a specific website.

12.1.2 Remote System Gathers Info Securely About User's Current Setup

In this step, the remote system determines the current setup for the user. The current setup must be authenticated, to ensure that the user truly has the setup that is claimed. Traditionally, this has been by checking the existence of files, generating checksums from those files, or by getting a serial number from a hardware dongle, although these traditional methods have difficulties since they can be generated locally by “hacked” software.

The authenticated read protocol described in Section 8.3 on page 664 can be used to accomplish this step. The use of random numbers has the advantage that the local user cannot capture a successful transaction and play it back on another computer system to fool the remote system.

12.1.3 Remote System Gives User Choice of Upgrade Possibilities & User Chooses

If there is more than one upgrade possibility, the various upgrade options are now presented to the user. The upgrade options could vary based on a number of factors, including, but not limited to:

-   -   current user setup     -   user's preference for payment schemes (e.g. single payment vs.         multiple payment)     -   number of other products owned by user

The user selects an appropriate upgrade and pays if necessary (by some scheme such as via a secure web site). What is important to note here is that the user chooses a specific upgrade and commences the upgrade operation.

12.1.4 Remote System Sends Upgrade Request to Local System

The remote system now instructs the local system to perform the upgrade. However, the local system can only accept an upgrade from the remote system if the remote system is also authenticated. This is effectively an authenticated write. The use of R_(U) in the signature prevents the upgrade message from being replayed on another ChipU.

If multiple keys are used, and each chip has a unique key, the remote system can use a serial number obtained from the current setup (authenticated by a common key) to lookup the unique key for use in the upgrade. Although the random number provides time varying messages, use of an unknown K that is different for each chip means that collection and examination of messages and their signatures is made even more difficult.

12.2 OEM Upgrades

OEM upgrades are effectively the same as remote upgrades, except that the user interacts with an OEM server for upgrade selection. The OEM server may send sub-requests to the manufacturer's remote server to provide authentication, upgrade availability lists, and base-level pricing information.

An additional level of authentication may be incorporated into the protocol to ensure that upgrade requests are coming from the OEM server, and not from a 3rd party. This can readily be incorporated into both authentication steps.

13 Choice of Signature Function

Given that all protocols make use of keyed signature functions, the choice of function is examined here.

Table 232 outlines the attributes of the applicable choices (see Section 5.2 on page 629 and Section 5.5 on page 636 for more information). The attributes are phrased so that the attribute is seen as an advantage.

TABLE 232 Attributes of Applicable Signature Functions HMAC- Triple Blow- Random HMAC- HMAC- RIPEM DES fish RC5 IDEA Sequences MD5 SHA1 D160 Free of patents • • • • • • Random key • • • generation Can be exported • • • • from the USA Fast • • • • Preferred Key 168¹ 128 128 128 512 128 160 160 Size (bits) for use in this application Block size (bits) 64 64 64 64 256 512 512 512 Cryptanalysis • • • • • Attack-Free (apart from weak keys) Output size given ≧N ≧N ≧N ≧N 128 128 160 160 input size N Low storage • • • • requirements Low silicon • • • • complexity NSA designed • • ¹Only gives protection equivalent to 112-bit DES

An examination of Table 232 shows that the choice is effectively between the 3 HMAC constructs and the Random Sequence. The problem of key size and key generation eliminates the Random Sequence. Given that a number of attacks have already been carried out on MD5 and since the hash result is only 128 bits, HMAC-MD5 is also eliminated. The choice is therefore between HMAC-SHA1 and HMAC-RIPEMD160. Of these, SHA-1 is the preferred function, since:

-   -   SHA-1 has been more extensively cryptanalyzed without being         broken;     -   SHA-1 requires slightly less intermediate storage than         RIPE-MD-160;     -   SHA-1 is algorithmically less complex than RIPE-MD-160;

Although SHA-1 is slightly faster than RIPE-MD-160, this was not a reason for choosing SHA-1.

13.1 HMAC-SHA1

The mechanism for authentication is the HMAC-SHA1 algorithm. This section examines the HMAC-SHA1 algorithm in greater detail than covered so far, and describes an optimization of the algorithm that requires fewer memory resources than the original definition.

13.1.1 HMAC

Given the following definitions:

-   -   H=the hash function (e.g. MD5 or SHA-1)     -   n=number of bits output from H (e.g. 160 for SHA-1, 128 bits for         MD5)     -   M=the data to which the MAC function is to be applied     -   K=the secret key shared by the two parties     -   ipad=0x36 repeated 64 times     -   opad=0x5C repeated 64 times

The HMAC algorithm is as follows:

-   a. Extend K to 64 bytes by appending 0×00 bytes to the end of K -   b. XOR the 64 byte string created in (1) with ipad -   c. append data stream M to the 64 byte string created in (2) -   d. Apply H to the stream generated in (3) -   e. XOR the 64 byte string created in (1) with opad -   f. Append the H result from (4) to the 64 byte string resulting from     (5) -   g. Apply H to the output of (6) and output the result

Thus: HMAC[M]=H[(K ⊕ opad)|H[(K ⊕ ipad)|M]]

-   -   The HMAC-SHA1 algorithm is simply HMAC with H=SHA-1.         13.1.2 SHA-1

The SHA1 hashing algorithm is described in the context of other hashing algorithms in Section 5.5.3.3 on page 640, and completely defined in [28]. The algorithm is summarized here.

Nine 32-bit constants are defined in Table 233. There are 5 constants used to initialize the chaining variables, and there are 4 additive constants.

TABLE 233 Constants used in SHA-1 Initial Chaining Values Additive Constants h₁ 0x67452301 y₁ 0x5A827999 h₂ 0xEFCDAB89 y₂ 0x6ED9EBA1 h₃ 0x98BADCFE y₃ 0x8F1BBCDC h₄ 0x10325476 y₄ 0xCA62C1D6 h₅ 0xC3D2E1F0

Non-optimized SHA-1 requires a total of 2912 bits of data storage:

-   -   Five 32-bit chaining variables are defined: H₁, H₂, H₃, H₄ and         H₅.     -   Five 32-bit working variables are defined: A, B, C, D, and E.     -   One 32-bit temporary variable is defined: t.     -   Eighty 32-bit temporary registers are defined: X₀₋₇₉.

The following functions are defined for SHA-1:

TABLE 234 Functions used in SHA-1 Symbolic Nomenclature Description + Addition modulo 2³² X

Y Result of rotating X left through Y bit positions f(X, Y, Z) (X

Y)

(

X

Z) g(X, Y, Z) (X

Y)

(X

Z)

(Y

Z) h(X, Y, Z) X ⊕ Y ⊕ Z

The hashing algorithm consists of firstly padding the input message to be a multiple of 512 bits and initializing the chaining variables H₁₋₅ with h₁₋₅. The padded message is then processed in 512-bit chunks, with the output hash value being the final 160-bit value given by the concatenation of the chaining variables: H₁|H₂|H₃|H₄|H₅.

The steps of the SHA-1 algorithm are now examined in greater detail.

13.1.2.1 Step 1. Preprocessing

The first step of SHA-1 is to pad the input message to be a multiple of 512 bits as follows and to initialize the chaining variables.

TABLE 235 Steps to follow to preprocess the input message Pad the input message Append a 1 bit to the message Append 0 bits such that the length of the padded message is 64-bits short of a multiple of 512 bits. Append a 64-bit value containing the length in bits of the original input message. Store the length as most significant bit through to least significant bit. Initialize the chaining H₁

h₁, H₂

h₂, H₃

h₃, H₄

h₄, variables H₅

h₅ 13.1.2.2 Step 2. Processing

The padded input message is processed in 512-bit blocks. Each 512-bit block is in the form of 16 32-bit words, referred to as InputWord₀₋₁₅.

TABLE 236 Steps to follow for each 512 bit block (InputWord₀₋₁₅) Copy the 512 For j = 0 to 15 input bits X_(j) = InputWord_(j) into X₀₋₁₅ Expand X₀₋₁₅ For j = 16 to 79 into X₁₆₋₇₉ X_(j)

((X_(j−3) ⊕ X_(j−8) ⊕ X_(j−14) ⊕ X_(j−16))

1) Initialize A

H₁, B

H₂, C

H₃, D

H₄, working E

H₅ variables Round 1 For j = 0 to 19 t

((A

5) + f(B, C, D) + E + Xj + y₁) E

D, D

C, C

(B

30), B

A, A

t Round 2 For j = 20 to 39 t

((A

5) + h(B, C, D) + E + Xj + y₂) E

D, D

C, C

(B

30), B

A, A

t Round 3 For j = 40 to 59 t

((A

5) + g(B, C, D) + E + Xj + y₃) E

D, D

C, C

(B

30), B

A, A

t Round 4 For j = 60 to 79 t

((A

5) + h(B, C, D) + E + Xj + y₄) E

D, D

C, C

(B

30), B

A, A

t Update chaining H₁

H₁ + A, H₂

H₂ + B, variables H₃

H₃ + C, H₄

H₄ + D, H₅

H₅ + E

The bold text is to emphasize the differences between each round.

13.1.2.3 Step 3. Completion

After all the 512-bit blocks of the padded input message have been processed, the output hash value is the final 160-bit value given by: H₁|H₂|H₃|H₄|H₅.

13.1.2.4 Optimization for Hardware Implementation

The SHA-1 Step 2 procedure is not optimized for hardware. In particular, the 80 temporary 32-bit registers use up valuable silicon on a hardware implementation. This section describes an optimization to the SHA-1 algorithm that only uses 16 temporary registers. The reduction in silicon is from 2560 bits down to 512 bits, a saving of over 2000 bits. It may not be important in some applications, but in the QA Chip storage space must be reduced where possible.

The optimization is based on the fact that although the original 16-word message block is expanded into an 80-word message block, the 80 words are not updated during the algorithm. In addition, the words rely on the previous 16 words only, and hence the expanded words can be calculated on-the-fly during processing, as long as we keep 16 words for the backward references. We require rotating counters to keep track of which register we are up to using, but the effect is to save a large amount of storage.

Rather than index X by a single value j, we use a 5 bit counter to count through the iterations. This can be achieved by initializing a 5-bit register with either 16 or 20, and decrementing it until it reaches 0. In order to update the 16 temporary variables as if they were 80, we require 4 indexes, each a 4-bit register. All 4 indexes increment (with wraparound) during the course of the algorithm.

TABLE 237 Optimised Steps to follow for each 512 bit block (InputWord₀₋₁₅) Initialize working A

H₁, B

H₂, C

H₃, D

H₄, variables E

H₅ N₁

13, N₂

8, N₃

2, N₄

0 Round 0 Do 16 times Copy the 512 X_(N4) = InputWord_(N4) input bits into [

N₁,

N₂,

N₃]_(optional)

N₄ X₀₋₁₅ Round 1A Do 16 times t

((A

5) + f(B, C, D) + E + X_(N4) + y₁) [

N₁,

N₂,

N₃]_(optional)

N₄ E

D, D

C, C

(B

30), B

A, A

t Round 1B Do 4 times X_(N4)

((X_(N1) ⊕ X_(N2) ⊕ X_(N3) ⊕ X_(N4))

1) t

((A

5) + f(B, C, D) + E + X_(N4) + y₁)

N₁,

N₂,

N₃,

N₄ E

D, D

C, C

(B

30), B

A, A

t Round 2 Do 20 times X_(N4)

((X_(N1) ⊕ X_(N2) ⊕ X_(N3) ⊕ X_(N4))

1) t

((A

5) + h(B, C, D) + E + X_(N4) + y₂)

N₁,

N₂,

N₃,

N₄ E

D, D

C, C

(B

30), B

A, A

t Round 3 Do 20 times X_(N4)

((X_(N1) ⊕ X_(N2) ⊕ X_(n3) ⊕ X_(N4))

1) t

((A

5) + g(B, C, D) + E + X_(N4) + y₃)

N₁,

N₂,

N₃,

N₄ E

D, D

C, C

(B

30), B

A, A

t Round 4 Do 20 times X_(N4)

((X_(N1) ⊕ X_(N2) ⊕ X_(N3) ⊕ X_(N4))

1) t

((A

5) + h(B, C, D) + E + X_(N4) + y₄)

N₁,

N₂,

N₃,

N₄ E

D, D

C, C

(B

30), B

A, A

t Update chaining H₁

H₁ + A, H₂

H₂ + B, variables H₃

H₃ + C, H₄

H₄ + D, H₅

H₅ + E

The bold text is to emphasize the differences between each round.

The incrementing of N₁, N₂, and N₃ during Rounds 0 and 1A is optional. A software implementation would not increment them, since it takes time, and at the end of the 16 times through the loop, all 4 counters will be their original values. Designers of hardware may wish to increment all 4 counters together to save on control logic.

Round 0 can be completely omitted if the caller loads the 512 bits of X₀₋₁₅.

14 Holding Out Against Attacks

The authentication protocols described in Section 7 on page 661 onward should be resistant to defeat by logical means. This section details each type of attack in turn with reference to the Read Authentication protocol.

14.1 Brute Force Attack

A brute force attack is guaranteed to break any protocol. However the length of the key means that the time for an attacker to perform a brute force attack is too long to be worth the effort.

An attacker only needs to break K to build a clone authentication chip. A brute force attack on K must therefore break a 160-bit key.

An attack against K requires a maximum of 2¹⁶⁰ attempts, with a 50% chance of finding the key after only 2¹⁵⁹ attempts. Assuming an array of a trillion processors, each running one million tests per second, 2¹⁵⁹ (7.3×10⁴⁷) tests takes 2.3×10²² years, which is longer than the total lifetime of the universe. There are around 100 million personal computers in the world. Even if these were all connected in an attack (e.g. via the Internet), this number is still 10,000 times smaller than the trillion-processor attack described. Further, if the manufacture of one trillion processors becomes a possibility in the age of nanocomputers, the time taken to obtain the key is still longer than the total lifetime of the universe.

14.2 Guessing the Key Attack

It is theoretically possible that an attacker can simply “guess the key”. In fact, given enough time, and trying every possible number, an attacker will obtain the key. This is identical to the brute force attack described above, where 2¹⁵⁹ attempts must be made before a 50% chance of success is obtained.

The chances of someone simply guessing the key on the first try is 2¹⁶⁰. For comparison, the chance of someone winning the top prize in a U.S. state lottery and being killed by lightning in the same day is only 1 in 2⁶¹ (78]. The chance of someone guessing the authentication chip key on the first go is 1 in 2¹⁶⁰, which is comparable to two people choosing exactly the same atoms from a choice of all the atoms in the Earth i.e. extremely unlikely.

14.3 Quantum Computer Attack

To break K, a quantum computer containing 160 qubits embedded in an appropriate algorithm must be built. As described in Section 5.7.1.7 on page 648, an attack against a 160-bit key is not feasible. An outside estimate of the possibility of quantum computers is that 50 qubits may be achievable within 50 years. Even using a 50 qubit quantum computer, 2¹¹⁰ tests are required to crack a 160 bit key. Assuming an array of 1 billion 50 qubit quantum computers, each able to try 2⁵⁰ keys in 1 microsecond (beyond the current wildest estimates) finding the key would take an average of 18 billion years.

14.4 Ciphertext Only Attack

An attacker can launch a ciphertext only attack on K by monitoring calls to Random and Read.

However, given that all these calls also reveal the plaintext as well as the hashed form of the plaintext, the attack would be transformed into a stronger form of attack—a known plaintext attack.

14.5 Known Plaintext Attack

It is easy to connect a logic analyzer to the connection between the System and the authentication chip, and thereby monitor the flow of data. This flow of data results in known plaintext and the hashed form of the plaintext, which can therefore be used to launch a known plaintext attack against K.

To launch an attack against K, multiple calls to Random and Test must be made (with the call to Test being successful, and therefore requiring a call to Read on a valid chip). This is straightforward, requiring the attacker to have both a system authentication chip and a consumable authentication chip. For each set of calls, an X, S_(K)[X] pair is revealed. The attacker must collect these pairs for further analysis.

The question arises of how many pairs must be collected for a meaningful attack to be launched with this data. An example of an attack that requires collection of data for statistical analysis is differential cryptanalysis (see Section 14.13 on page 703). However, there are no known attacks against SHA-1 or HMAC-SHA1 [7][7][7], so there is no use for the collected data at this time.

14.6 Chosen Plaintext Attacks

The golden rule for the QA Chip is that it never signs something that is simply given to it—i.e. it never lets the user choose the message that is signed.

Although the attacker can choose both R_(T) and possibly M, ChipA advances its random number R_(A) with each call to Read. The resultant message X therefore contains 160 bits of changing data each call that are not chosen by the attacker.

To launch a chosen text attack the attacker would need to locate a chip whose R was the desired R. This makes the search effectively impossible.

14.7 Adaptive Chosen Plaintext Attacks

The HMAC construct provides security against all forms of chosen plaintext attacks [7]. This is primarily because the HMAC construct has 2 secret input variables (the result of the original hash, and the secret key). Thus finding collisions in the hash function itself when the input variable is secret is even harder than finding collisions in the plain hash function. This is because the former requires direct access to SHA-1 in order to generate pairs of input/output from SHA-1.

Since R changes with each call to Read, the user cannot choose the complete message. The only value that can be collected by an attacker is HMAC[R₁|R₂|M₂]. These are not attacks against the SHA-1 hash function itself, and reduce the attack to a differential cryptanalysis attack (see Section 14.13 on page 703), examining statistical differences between collected data. Given that there is no differential cryptanalysis attack known against SHA-1 or HMAC, the protocols are resistant to the adaptive chosen plaintext attacks.

14.8 Purposeful Error Attack

An attacker can only launch a purposeful error attack on the Test function, since this is the only function in the Read protocol that validates input against the keys.

With the Test function, a 0 value is produced if an error is found in the input—no further information is given. In addition, the time taken to produce the 0 result is independent of the input, giving the attacker no information about which bit(s) were wrong.

A purposeful error attack is therefore fruitless.

14.9 Chaining Attack

Any form of chaining attack assumes that the message to be hashed is over several blocks, or the input variables can somehow be set. The HMAC-SHA1 algorithm used by Protocol C1 only ever hashes one or two 512-bit blocks. Chaining attacks are not possible when only one block is used, and are extremely limited when two blocks are used.

14.10 Birthday Attack

The strongest attack known against HMAC is the birthday attack, based on the frequency of collisions for the hash function [7][7]. However this is totally impractical for minimally reasonable hash functions such as SHA-1. And the birthday attack is only possible when the attacker has control over the message that is hashed.

Since in the protocols described for the QA Chip, the message to be signed is never chosen by the attacker (at least one 160-bit R value is chosen by the chip doing the signing), the attacker has no control over the message that is hashed. An attacker must instead search for a collision message that hashes to the same value (analogous to finding one person who shares your birthday).

The clone chip must therefore attempt to find a new value R₂ such that the hash of R₁, R₂ and a chosen M₂ yields the same hash value as H[R₁|R₂|M]. However ChipT does not reveal the correct hash value (the Test function only returns 1 or 0 depending on whether the hash value is correct).

Therefore the only way of finding out the correct hash value (in order to find a collision) is to interrogate a real ChipA. But to find the correct value means to update M, and since the decrement-only parts of M are one-way, and the read-only parts of M cannot be changed, a clone consumable would have to update a real consumable before attempting to find a collision. The alternative is a brute force attack search on the Test function to find a success (requiring each clone consumable to have access to a System consumable). A brute force search, as described above, takes longer than the lifetime of the universe, in this case, per authentication.

There is no point for a clone consumable to launch this kind of attack.

14.11 Substitution With A Complete Lookup Table

The random number seed in each System is 160 bits. The best case situation for an attacker is that no state data has been changed. Assuming also that the clone consumable does not advance its R, there is a constant value returned as M. A clone chip must therefore return S_(K)[R|c] (where c is a 10 constant), which is a 160 bit value.

Assuming a 160-bit lookup of a 160-bit result, this requires 2.9×10⁴⁹ bytes, or 2.6×10³⁷ terabytes, certainly more space than is feasible for the near future. This of course does not even take into account the method of collecting the values for the ROM. A complete lookup table is therefore completely impossible.

14.12 Substitution With A Sparse Lookup Table

A sparse lookup table is only feasible if the messages sent to the authentication chip are somehow predictable, rather than effectively random.

The random number R is seeded with an unknown random number, gathered from a naturally System authentication chip's Random function, and iterating some random event. There is no possibility for a clone manufacturer to know what the possible range of R is for all Systems, since each bit has an unrelated chance of being 1 or 0.

Since the range of R in all systems is unknown, it is not possible to build a sparse lookup table that can be used in all systems. The general sparse lookup table is therefore not a possible attack.

However, it is possible for a clone manufacturer to know what the range of R is for a given System. This can be accomplished by loading a LFSR with the current result from a call to a specific number of times into the future. If this is done, a special ROM can be built which will only contain the responses for that particular range of R, i.e. a ROM specifically for the consumables of that particular System. But the attacker still needs to place correct information in the ROM. The attacker will therefore need to find a valid authentication chip and call it for each of the values in R. Suppose the clone authentication chip reports a full consumable, and then allows a single use before simulating loss of connection and insertion of a new full consumable. The clone consumable would therefore need to contain responses for authentication of a full consumable and authentication of a partially used consumable. The worst case ROM contains entries for full and partially used consumables for R over the lifetime of System. However, a valid authentication chip must be used to generate the information, and be partially used in the process. If a given System only produces n R-values, the sparse lookup-ROM required is 20n bytes (20=160/8) multiplied by the number of different values for M. The time taken to build the ROM depends on the amount of time enforced between calls to Read.

After all this, the clone manufacturer must rely on the consumer returning for a refill, since the cost of building the ROM in the first place consumes a single consumable. The clone manufacturer's business in such a situation is consequently in the refills.

The time and cost then, depends on the size of R and the number of different values for M that must be incorporated in the lookup. In addition, a custom clone consumable ROM must be built to match each and every System, and a different valid authentication chip must be used for each System (in order to provide the full and partially used data). The use of an authentication chip in a System must therefore be examined to determine whether or not this kind of attack is worthwhile for a clone manufacturer.

As an example, of a camera system that has about 10,000 prints in its lifetime. Assume it has a single Decrement Only value (number of prints remaining), and a delay of 1 second between calls to Read. In such a system, the sparse table will take about 3 hours to build, and consumes 100K. Remember that the construction of the ROM requires the consumption of a valid authentication chip, so any money charged must be worth more than a single consumable and the clone consumable combined. Thus it is not cost effective to perform this function for a single consumable (unless the clone consumable somehow contained the equivalent of multiple authentic consumables).

If a clone manufacturer is going to go to the trouble of building a custom ROM for each owner of a System, an easier approach would be to update System to completely ignore the authentication chip.

Consequently, this attack is possible as a per-System attack, and a decision must be made about the chance of this occurring for a given System/Consumable combination. The chance will depend on the cost of the consumable and authentication chips, the longevity of the consumable, the profit margin on the consumable, the time taken to generate the ROM, the size of the resultant ROM, and whether customers will come back to the clone manufacturer for refills that use the same clone chip etc.

14.13 Differential Cryptanalysis

Existing differential attacks are heavily dependent on the structure of S boxes, as used in DES and other similar algorithms. Although HMAC-SHA1 has no S boxes, an attacker can undertake a differential-like attack by undertaking statistical analysis of:

-   -   Minimal-difference inputs, and their corresponding outputs     -   Minimal-difference outputs, and their corresponding inputs

To launch an attack of this nature, sets of input/output pairs must be collected. The collection can be via known plaintext, or from a partially adaptive chosen plaintext attack. Obviously the latter, being chosen, will be more useful.

Hashing algorithms in general are designed to be resistant to differential analysis. SHA-1 in particular has been specifically strengthened, especially by the 80 word expansion so that minimal differences in input will still produce outputs that vary in a larger number of bit positions (compared to 128 bit hash functions). In addition, the information collected is not a direct SHA-1 input/output set, due to the nature of the HMAC algorithm. The HMAC algorithm hashes a known value with an unknown value (the key), and the result of this hash is then rehashed with a separate unknown value. Since the attacker does not know the secret value, nor the result of the first hash, the inputs and outputs from SHA-1 are not known, making any differential attack extremely difficult.

There are no known differential attacks against SHA-1 or HMAC-SHA-1 [56][56].

The following is a more detailed discussion of minimally different inputs and outputs from the QA Chip.

14.13.1 Minimal Difference InpuTS

This is where an attacker takes a set of X, S_(K)[X] values where the X values are minimally different, and examines the statistical differences between the outputs S_(K)[X]. The attack relies on X values that only differ by a minimal number of bits. The question then arises as to how to obtain minimally different X values in order to compare the S_(K)[X] values.

Although the attacker can choose both R_(T) and possibly M, ChipA advances its random number R_(A) with each call to Read. The resultant X therefore contains 160 bits of changing data each call, and is therefore not minimally different.

14.13.2 Minimal Difference Outputs

This is where an attacker takes a set of X, S_(K)[X] values where the S_(K)[X] values are minimally different, and examines the statistical differences between the X values. The attack relies on S_(K)[X] values that only differ by a minimal number of bits.

There is no way for an attacker to generate an X value for a given S_(K)[X]. To do so would violate the fact that S is a one-way function (HMAC-SHA1). Consequently the only way for an attacker to mount an attack of this nature is to record all observed X, S_(K)[X] pairs in a table. A search must then be made through the observed values for enough minimally different S_(K)[X] values to undertake a statistical analysis of the X values.

14.14 Message Substitution Attacks

In order for this kind of attack to be carried out, a clone consumable must contain a real authentication chip, but one that is effectively reusable since it never gets decremented. The clone authentication chip would intercept messages, and substitute its own. However this attack does not give success to the attacker.

A clone authentication chip may choose not to pass on a Write command to the real authentication chip. However the subsequent Read command must return the correct response (as if the Write had succeeded). To return the correct response, the hash value must be known for the specific R and M. An attacker can only determine the hash value by actually updating M in a real Chip, which the attacker does not want to do. Even changing the R sent by System does not help since the System authentication chip must match the R during a subsequent Test.

A message substitution attack would therefore be unsuccessful. This is only true if System updates the amount of consumable remaining before it is used.

14.15 Reverse Engineering The Key Generator

If a pseudo-random number generator is used to generate keys, there is the potential for a clone manufacture to obtain the generator program or to deduce the random seed used. This was the way in which the security layer of the Netscape browser was initially broken [33].

14.16 Bypassing The Authentication Process

The System should ideally update the consumable state data before the consumable is used, and follow every write by a read (to authenticate the write). Thus each use of the consumable requires an authentication. If the System adheres to these two simple rules, a clone manufacturer will have to simulate authentication via a method above (such as sparse ROM lookup).

14.17 Reuse Of Authentication Chips

Each use of the consumable requires an authentication. If a consumable has been used up, then its authentication chip will have had the appropriate state-data values decremented to 0. The chip can therefore not be used in another consumable.

Note that this only holds true for authentication chips that hold Decrement-Only data items. If there is no state data decremented with each usage, there is nothing stopping the reuse of the chip. This is the basic difference between Presence-Only authentication and Consumable Lifetime authentication. All described protocols allow both.

The bottom line is that if a consumable has Decrement Only data items that are used by the System, the authentication chip cannot be reused without being completely reprogrammed by a valid programming station that has knowledge of the secret key (e.g. an authorized refill station).

14.18 Management Decision To Omit Authentication To Save Costs

Although not strictly an external attack, a decision to omit authentication in future Systems in order to save costs will have widely varying effects on different markets.

In the case of high volume consumables, it is essential to remember that it is very difficult to introduce authentication after the market has started, as systems requiring authenticated consumables will not work with older consumables still in circulation. Likewise, it is impractical to discontinue authentication at any stage, as older Systems will not work with the new, unauthenticated, consumables. In the second case, older Systems can be individually altered by replacing the System program code.

Without any form of protection, illegal cloning of high volume consumables is almost certain. However, with the patent and copyright protection, the probability of illegal cloning may be, say 50%. However, this is not the only loss possible. If a clone manufacturer were to introduce clone consumables which caused damage to the System (e.g. clogged nozzles in a printer due to poor quality ink), then the loss in market acceptance, and the expense of warranty repairs, may be significant.

In the case of a specialized pairing, such as a car/car-keys, or door/door-key, or some other similar situation, the omission of authentication in future systems is trivial and without repercussions. This is because the consumer is sold the entire set of System and Consumable authentication chips at the one time.

14.19 Garrote/Bribe Attack

If humans do not know the key, there is no amount of force or bribery that can reveal them. The use of ChipF and the ReplaceKey protocol is specifically designed to avoid the requirement of the programming station having to know the new key. However ChipF must be told the new key at some stage, and therefore it is the person(s) who enter the new key into ChipF that are at risk.

The level of security against this kind of attack is ultimately a decision for the System/Consumable owner, to be made according to the desired level of service.

For example, a car company may wish to keep a record of all keys manufactured, so that a person can request a new key to be made for their car. However this allows the potential compromise of the entire key database, allowing an attacker to make keys for any of the manufacturer's existing cars. It does not allow an attacker to make keys for any new cars. Of course, the key database itself may also be encrypted with a further key that requires a certain number of people to combine their key portions together for access. If no record is kept of which key is used in a particular car, there is no way to make additional keys should one become lost. Thus an owner will have to replace his car's authentication chip and all his car-keys. This is not necessarily a bad situation.

By contrast, in a consumable such as a printer ink cartridge, the one key combination is used for all Systems and all consumables. Certainly if no backup of the keys is kept, there is no human with knowledge of the key, and therefore no attack is possible. However, a no-backup situation is not desirable for a consumable such as ink cartridges, since if the key is lost no more consumables can be made. The manufacturer should therefore keep a backup of the key information in several parts, where a certain number of people must together combine their portions to reveal the full key information. This may be required if case the chip programming station needs to be reloaded.

In any case, none of these attacks are against the authenticated read protocol, since no humans are involved in the authentication process.

Logical Interface

15 Introduction

The QA Chip has a physical and a logical external interface. The physical interface defines how the QA Chip can be connected to a physical System, while the logical interface determines how that System can communicate with the QA Chip. This section deals with the logical interface.

15.1 Operating Modes

The QA Chip has four operating modes—Idle Mode, Program Mode, Trim Mode and Active Mode.

-   -   Idle Mode is used to allow the chip to wait for the next         instruction from the System.     -   Trim Mode is used to determine the clock speed of the chip and         to trim the frequency during the initial programming stage of         the chip (when Flash memory is garbage). The clock frequency         must be trimmed via Trim Mode before Program Mode is used to         store the program code.     -   Program Mode is used to load up the operating program code, and         is required because the operating program code is stored in         Flash memory instead of ROM (for security reasons).     -   Active Mode is used to execute the specific authentication         command specified by the System. Program code is executed in         Active Mode. When the results of the command have been returned         to the System, the chip enters Idle Mode to wait for the next         instruction.         15.1.1 Idle Mode

The QA Chip starts up in Idle Mode. When the Chip is in Idle Mode, it waits for a command from the master by watching the primary id on the serial line.

-   -   If the primary id matches the global id (0x00, common to all QA         Chips), and the following byte from the master is the Trim Mode         id byte, the QA Chip enters Trim Mode and starts counting the         number of internal clock cycles until the next byte is received.     -   If the primary id matches the global id (0x00, common to all QA         Chips), and the following byte from the master is the Program         Mode id byte, the QA Chip enters Program Mode.     -   If the primary id matches the global id (0x00, common to all QA         Chips), and the following byte from the master is the Active         Mode id byte, the QA Chip enters Active Mode and executes         startup code, allowing the chip to set itself into a state to         receive authentication commands (includes setting a local ID).     -   If the primary id matches the chip's local ID, and the following         byte is a valid command code, the QA Chip enters Active Mode,         allowing the command to be executed.

The valid 8-bit serial mode values sent after a global id are as shown in Table 238. They are specified to minimize the chances of them occurring by error after a global id (e.g. 0xFF and 0x00 are not used):

TABLE 238 Id byte values to place chip in specific mode Value Interpretation 10100101 (0xA5) Trim Mode 10001110 (0x8E) Program Mode 01111000 (0x78) Active Mode 15.1.2 Trim Mode

Trim Mode is enabled by sending a global id byte (0x00) followed by the Trim Mode command byte. The purpose of Trim Mode is to set the trim value (an internal register setting) of the internal ring oscillator so that Flash erasures and writes are of the correct duration. This is necessary due to the variation of the clock speed due to process variations. If writes an erasures are too long, the Flash memory will wear out faster than desired, and in some cases can even be damaged.

Trim Mode works by measuring the number of system clock cycles that occur inside the chip from the receipt of the Trim Mode command byte until the receipt of a data byte. When the data byte is received, the data byte is copied to the trim register and the current value of the count is transmitted to the outside world.

Once the count has been transmitted, the QA Chip returns to Idle Mode.

At reset, the internal trim register setting is set to a known value r. The external user can now perform the following operations:

-   -   send the global id+write followed by the Trim Mode command byte     -   send the 8-bit value v over a specified time t     -   send a stop bit to signify no more data     -   send the global id+read followed by the Trim Mode command byte     -   receive the count c     -   send a stop bit to signify no more data

At the end of this procedure, the trim register will be v, and the external user will know the relationship between external time t and internal time c. Therefore a new value for v can be calculated.

The Trim Mode procedure can be repeated a number of times, varying both t and v in known ways, measuring the resultant c. At the end of the process, the final value for v is established (and stored in the trim register for subsequent use in Program Mode). This value v must also be written to the flash for later use (every time the chip is placed in Active Mode for the first time after power-up).

15.1.3 Program Mode

Program Mode is enabled by sending a global id byte (0x00) followed by the Program Mode command byte.

The QA Chip determines whether or not the internal fuse has been blown (by reading 32-bit word 0 of the information block of flash memory).

If the fuse has been blown the Program Mode command is ignored, and the QA Chip returns to Idle Mode.

If the fuse is still intact, the chip enters Program Mode and erases the entire contents of Flash memory. The QA Chip then validates the erasure. If the erasure was successful, the QA Chip receives up to 4096 bytes of data corresponding to the new program code and variable data. The bytes are transferred in order byte₀ to byte₄₀₉₅.

Once all bytes of data have been loaded into Flash, the QA Chip returns to Idle Mode.

Note that Trim Mode functionality must be performed before a chip enters Program Mode for the first time.

Once the desired number of bytes have been downloaded in Program Mode, the LSS Master must wait for 80 μs (the time taken to write two bytes to flash at nybble rates) before sending the new transaction (eg Active Mode). Otherwise the last nybbles may not be written to flash.

15.1.4 Active Mode

Active Mode is entered either by receiving a global id byte (0x00) followed by the Active Mode command byte, or by sending a local id byte followed by a command opcode byte and an appropriate number of data bytes representing the required input parameters for that opcode.

In both cases, Active Mode causes execution of program code previously stored in the flash memory via Program Mode. As a result, we never enter Active Mode after Trim Mode, without a Program Mode in between. However once programmed via Program Mode, a chip is allowed to enter Active Mode after power-up, since valid data will be in flash.

If Active Mode is entered by the global id mechanism, the QA Chip executes specific reset startup code, typically setting up the local id and other IO specific data.

If Active Mode is entered by the local id mechanism, the QA Chip executes specific code depending on the following byte, which functions as an opcode. The opcode command byte format is shown in Table 239:

TABLE 239 Command byte bits Description 2-0 Opcode 5-3

opcode 7-6 count of number of bits set in opcode (0 to 3)

The interpretation of the 3-bit opcode is shown in Table 240:

TABLE 240 QA Chip opcodes Op² Mn³ Description 000 RST Reset 001 RND Random 010 RDM Read M 011 TST Test 100 WRM Write M with no authentication 101 WRA Write with Authentication (to M, P, or K) 110 chip specific - reserved for ChipF, ChipS etc 111 chip specific - reserved for ChipF, ChipS etc ²Opcode ³Mnemonic

The command byte is designed to ensure that errors in transmission are detected. Regular QA Chip commands are therefore comprised of an opcode plus any associated parameters. The commands are listed in Table 241:

TABLE 241 QA Chip commands Input Output Command opcode Additional parms Return value Reset RST — — Random RND — [20] Read RDM [1, 1, 20] [20, 64, 20]⁴ Test TST [1, 20, 64, 20] 89⁵ if successful, 76 if not Write WRM [1, 64, 20] 89 if successful, 76 if not WriteAuth WRA 76 [20, 64, 20] 89 if successful, 76 if not ReplaceKey WRA 89 76 [1, 20, 20, 20] 89 if successful, 76 if not SetPermissions WRA 89 89 [1, 1, 20, 4, 20] [4] SignM⁶ ChipS [1, 20, 20, 64, 20, 64] [20, 64, 20] only SignP⁷ ChipS [1, 20, 20, 4, 20, 4] [20, 64, 20] only GetProgKey ChipF [1, 20] [20, 20, 20] only SetPartialKey ChipF [1, 4] 89 if only successful, 76 if not ⁴[n, m] = list of parameters where n bytes for first parameter, and m bytes for the second etc. ⁵n = actual byte pattern required (in hex). The bytes 0x76 and 0x89 were chosen as the boolean values 0 and 1 as they are inverses of each other, and should not be generated accidentally. ⁶It is expected that most QA Chips will implement SignM as a function that returns 0x00. Only a limited number of chips will be programmed to allow SignM functionality. It is included here as an example of how signatures can be generated for authenticated writes. ⁷It is expected that most QA Chips will implement SignP as a function that returns 0x00. Only a limited number of chips will be programmed to allow SignP functionality. It is included here as an example of how signatures can be generated for authenticated writes.

Apart from the Reset command, the next four commands are the commands most likely to be used during regular operation. The next three commands are used to provide authenticated writes (which are expected to be uncommon). The final set of commands (including SignM), are expected to be specially implemented on ChipS and ChipF QA Chips only.

The input parameters are sent in the specified order, with each parameter being sent least significant byte first and most significant byte last.

Return (output) values are read in the same way—least significant byte first and most significant byte last. The client must know how many bytes to retrieve. The QA Chip will time out and return to Idle Mode if an incorrect number of bytes is provided or read.

In most cases, the output bytes from one chip's command (the return values) can be fed directly as the input bytes to another chip's command. An example of this is the RND and RD commands. The output data from a call to RND on a trusted QA Chip does not have to be kept by the System. Instead, the System can transfer the output bytes directly to the input of the non-trusted QA Chip's RD command. The description of each command points out where this is so.

Each of the commands is examined in detail in the subsequent sections. Note that some algorithms are specifically designed because flash memory is assumed for the implementation of non-volatile variables.

15.1.5 Non Volatile Variables

The memory within the QA Chip contains some non-volatile (Flash) memory to store the variables required by the authentication protocol. Table 242 summarizes the variables.

TABLE 242 Non volatile variables required by the authentication protocol Name Size (bits) Description N 8 Number of keys known to the chip T 8 Number of vectors M is broken into K_(n) 160 per key, Array of N secret keys used for R_(K) 160 for R_(K) calculating F_(Kn)[X] where K_(n) is the nth element of the array. Each K_(n) must not be stored directly in the QA Chip. Instead, each chip needs to store a single random number R_(K) (different for each chip), K_(n)⊕R_(K), and

K_(n)⊕R_(K). The stored K_(n)⊕R_(K) can be XORed with R_(K) to obtain the real K_(n). Although

K_(n)⊕R_(K) must be stored to protect against differential attacks, it is not used. R 160 Current random number used to ensure time varying messages. Each chip instance must be seeded with a different initial value. Changes for each signature generation. M_(T) 512 per M Array of T memory vectors. Only M₀ can be written to with an authorized write, while all Ms can be written to in an unauthorized write. Writes to M₀ are optimized for Flash usage, while updates to any other M_(n) are expensive with regards to Flash utilization, and are expected to be only performed once per section of M_(n). M₁ contains T and N in ReadOnly form so users of the chip can know these two values. P_(T+N) 32 per P T + N element array of access permissions for each part of M. Entries n = {0 . . . T − 1} hold access permissions for non- authenticated writes to M_(n) (no key required). Entries n = {T to T + N − 1} hold access permissions for authenticated writes to M₀ for K_(n). Permission choices for each part of M are Read Only, Read/Write, and Decrement Only MinTicks 32 The minimum number of clock ticks between calls to key-based functions. Note that since these variables are in Flash memory, writes should be minimized. The it is not a simple matter to write a new value to replace the old. Care must be taken with flash endurance, and speed of access. This has an effect on the algorithms used to change Flash memory based registers. For example, Flash memory should not be used as a shift register. A reset of the QA Chip has no effect on the non-volatile variables. 15.1.5.1 M and P

M_(n) contains application specific state data, such as serial numbers, batch numbers, and amount of consumable remaining. M_(n) can be read using the Read command and written to via the Write and WriteA commands.

M₀ is expected to be updated frequently, while each part of M_(1-n) should only be written to once. Only M₀ can be written to via the WriteA command.

M₁ contains the operating parameters of the chip as shown in Table 243, and M_(2-n), are application specific.

TABLE 243 Interpretation of M₁ Length Bits interpretation 8 7-0 Number of available keys 8 15-8  Number of available M vectors 16 31-16 Revision of chip 96 127-32  Manufacture id information 128 255-128 Serial number 8 263-256 Local id of chip 248 511-264 reserved

Each M_(n) is 512 bits in length, and is interpreted as a set of 16×32-bit words. Although M_(n) may contain a number of different elements, each 32-bit word differs only in write permissions. Each 32-bit word can always be read. Once in client memory, the 512 bits can be interpreted in any way chosen by the client. The different write permissions for each P are outlined in Table 244:

TABLE 244 Write permissions Data type permission description Read Only Can never be written to ReadWrite Can always be written to Decrement Can only be written to if the new value is less Only than the old value. Decrement Only values can be any multiple of 32 bits.

To accomplish the protection required for writing, a 2-bit permission value P is defined for each of the 32-bit words. Table 245 defines the interpretation of the 2-bit permission bit-pattern:

TABLE 245 Permission bit interpretation Bits Op Interpretation Action taken during Write command 00 RW ReadWrite The new 32-bit value is always written to M[n]. 01 MSR Decrement The new 32-bit value is only Only (Most written to M[n] if it is Significant less than the value currently Region) in M[n]. This is used for access to the Most Significant 16 bits of a Decrement Only number. 10 NMSR Decrement The new 32-bit value is only Only (Not written to M[n] if M[n − 1] the Most could also be written. The Significant NMSR access mode allows multiple Region) precision values of 32 bits and more (multiples of 32 bits) to decrement. 11 RO Read Only The new 32-bit value is ignored. M[n] is left unchanged.

The 16 sets of permission bits for each 512 bits of M are gathered together in a single 32-bit variable P, where bits 2n and 2n+1 of P correspond to word n of M as follows:

Each 2-bit value is stored as a pair with the msb in bit 1, and the lsb in bit 0. Consequently, if words 0 to 5 of M had permission MSR, with words 6-15 of M permission RO, the 32-bit P variable would be 0xFFFFF555:

11-11-11-11-11-11-11-11-11-11-01-01-01-01-01-01

During execution of a Write and WriteA command, the appropriate Permissions[n] is examined for each M[n] starting from n=15 (msw of M) to n=0 (lsw of M), and a decision made as to whether the ne M[n] value will replace the old. Note that it is important to process the M[n] from msw to lsw to correctly interpret the access permissions.

Permissions are set and read using the QA Chip's SetPermissions command. The default for P is all 0s (RW) with the exception of certain parts of M₁.

Note that the Decrement Only comparison is unsigned, so any Decrement Only values that require negative ranges must be shifted into a positive range. For example, a consumable with a Decrement Only data item range of −50 to 50 must have the range shifted to be 0 to 100. The System must then interpret the range 0 to 100 as being −50 to 50. Note that most instances of Decrement Only ranges are N to 0, so there is no range shift required.

For Decrement Only data items, arrange the data in order from most significant to least significant 32-bit quantities from M[n] onward. The access mode for the most significant 32 bits (stored in M[n]) should be set to MSR. The remaining 32-bit entries for the data should have their permissions set to NMSR.

If erroneously set to NMSR, with no associated MSR region, each NMSR region will be considered independently instead of being a multi-precision comparison.

Examples of allocating M and Permission bits can be found in [86].

15.1.5.2 K and R_(K)

K is the 160-bit secret key used to protect M and to ensure that the contents of M are valid (when M is read from a non trusted chip). K is initially programmed after manufacture, and from that point on, K can only be updated to a new value if the old K is known. Since K must be kept secret, there is no command to directly read it.

K is used in the keyed one-way hash function HMAC-SHA1. As such it should be programmed with a physically generated random number, gathered from a physically random phenomenon. K must NOT be generated with a computer-run random number generator. The security of the QA Chips depends on K being generated in a way that is not deterministic.

Each K_(n) must not be stored directly in the QA Chip. Instead, each chip needs to store a single random number R_(K) (different for each chip), K_(n)⊕R_(K), and

K_(n)⊕R_(K). The stored K_(n)⊕R_(K) can be XORed with R_(K) to obtain the real K_(n). Although

K_(n)R_(K) must be stored to protect against differential attacks, it is not used.

15.1.5.3 R

R is a 160-bit random number seed that is set up after manufacture (when the chip is programmed) and from that point on, cannot be changed. R is used to ensure that each signed item contains time varying information (not chosen by an attacker), and each chip's R is unrelated from one chip to the next.

R is used during the Test command to ensure that the R from the previous call to Random was used as the session key in generating the signature during Read. Likewise, R is used during the WriteAuth command to ensure that the R from the previous call to Read was used as the session key during generation of the signature in the remote Authenticated chip.

The only invalid value for R is 0. This is because R is changed via a 160-bit maximal period LFSR (Linear Feedback Shift Register) with taps on bits 0, 2, 3, and 5, and is changed only by a successful call to a signature generating function (e.g. Test, WriteAuth).

The logical security of the QA Chip relies not only upon the randomness of K and the strength of the HMAC-SHA1 algorithm. To prevent an attacker from building a sparse lookup table, the security of the QA Chip also depends on the range of R over the lifetime of all Systems. What this means is that an attacker must not be able to deduce what values of R there are in produced and future Systems. Ideally, R should be programmed with a physically generated random number, gathered from a physically random phenomenon (must not be deterministic). R must NOT be generated with a computer-run random number generator.

15.1.5.4 MinTicks

There are two mechanisms for preventing an attacker from generating multiple calls to key-based functions in a short period of time. The first is an internal ring oscillator that is temperature-filtered. The second mechanism is the 32-bit MinTicks variable, which is used to specify the minimum number of QA Chip clock ticks that must elapse between calls to key-based functions.

The MinTicks variable is set to a fixed value when the QA Chip is programmed. It could possibly be stored in M₁.

The effective value of MinTicks depends on the operating clock speed and the notion of what constitutes a reasonable time between key-based function calls (application specific). The duration of a single tick depends on the operating clock speed. This is the fastest speed of the ring oscillator generated clock (i.e. at the lowest valid operating temperature).

Once the duration of a tick is known, the MinTicks value can to be set. The value for MinTicks will be the minimum number of ticks required to pass between calls to the key-based functions (there is no need to protect Random as this produces the same output each time it is called multiple times in a row). The value is a real-time number, and divided by the length of an operating tick.

It should be noted that the MinTicks variable only slows down an attacker and causes the attack to cost more since it does not stop an attacker using multiple System chips in parallel.

15.1.6 GetProgramKey

-   Input: n, R_(E)=[1 byte, 20 bytes] -   Output: R_(L), E_(Kx)[S_(Kn)[R_(E)|R_(L)|C₃]],     S_(Kx)[R_(L)|E_(Kx)[S_(Kn)[R_(E)|R_(L)|C₃]|C₃]=[20, 20, 20] -   Changes: R_(L)

Note: The GetProgramKey command is only implemented in ChipF, and not in all QA Chips. The GetProgramKey command is used to produce the bytestream required for updating a specified key in ChipP. Only an QA Chip programmed with the correct values of the old K_(n) can respond correctly to the GetProgramKey request. The output bytestream from the Random command can be fed as the input bytestream to the ReplaceKey command on the QA Chip being programmed (ChipP).

The input bytestream consists of the appropriate opcode followed by the desired key to generate the signature, followed by 20 bytes of R_(E) (representing the random number read in from ChipP).

The local random number R_(L) is advanced, and signed in combination with R_(E) and C₃ by the chosen key to generate a time varying secret number known to both ChipF and ChipP. This signature is then XORed with the new key K_(x) (this encrypts the new key). The first two output parameters are signed with the old key to ensure that ChipP knows it decoded K_(x) correctly.

This whole procedure should only be allowed a given number of times. The actual number can conveniently be stored in the local M₀[0] (eg word 0 of M₀) with ReadOnly permission. Of course another chip could perform an Authorised write to update the number (via a ChipS) should it be desired.

The GetProgramKey command is implemented by the following steps:

Loop through all of Flash, reading each word (will trigger checks) Accept n Restrict n to N Accept R_(E) If (M₀[0] = 0) Output 60 bytes of 0x00 # no more keys allowed to be generated from this chipF Done EndIf Advance R_(L) SIG

S_(Kn)[R_(L)|R_(E)|C₃] # calculation must take constant time Tmp

SIG ⊕ K_(X) Output R_(L) Output Tmp Decrement M₀[0] #  reduce the number of allowable key generations by 1 SIG

S_(KX)[R_(L)|Tmp|C₃] # calculation must take constant time Output SIG 15.1.7 Random

-   -   Input: None     -   Output: R_(L)=[20 bytes]     -   Changes: None

The Random command is used by a client to obtain an input for use in a subsequent authentication procedure. Since the Random command requires no input parameters, it is therefore simply 1 byte containing the RND opcode.

The output of the Random command from a trusted QA Chip can be fed straight into the non-trusted chip's Read command as part of the input parameters. There is no need for the client to store them at all, since they are not required again. However the Test command will only succeed if the data passed to the Read command was obtained first from the Random command.

If a caller only calls the Random function multiple times, the same output will be returned each time.

R will only advance to the next random number in the sequence after a successful call to a function that returns or tests a signature (e.g. Test, see Section 15.1.13 on page 725 for more information).

The Random command is implemented by the following steps:

-   -   Loop through all of Flash, reading each word (will trigger         checks)     -   Output R_(L)         15.1.8 Read     -   Input: n, t, R_(E)=[1 byte, 1 byte, 20 bytes]     -   Output: R_(L), M_(Lt), S_(Kn)[R_(E)|R_(L)|C₁|M_(Lt)]=[20 bytes,         64 bytes, 20 bytes]     -   Changes: R_(L)

The Read command is used to read the entire state data (M_(t)) from an QA Chip. Only an QA Chip programmed with the correct value of K_(n) can respond correctly to the Read request. The output bytestream from the Read command can be fed as the input bytestream to the Test command on a trusted QA Chip for verification, with M_(t) stored for later use if Test returns success.

The input bytestream consists of the RD opcode followed by the key number to use for the signature, which M to read, and the bytes 0-19 of R_(E). 23 bytes are transferred in total. R_(E) is obtained by calling the trusted QA Chip's Random command. The 20 bytes output by the trusted chip's Random command can therefore be fed directly into the non-trusted chip's Read command, with no need for these bits to be stored by System.

Calls to Read must wait for MinTicksRemaining to reach 0 to ensure that a minimum time will elapse between calls to Read.

The output values are calculated, MinTicksRemaining is updated, and the signature is returned. The contents of M_(Lt) are transferred least significant byte to most significant byte. The signature S_(Kn)[R_(E)|R_(L)|C₁|M_(Lt)] must be calculated in constant time.

The next random number is generated from R using a 160-bit maximal period LFSR (tap selections on bits 5, 3, 2, and 0). The initial 160-bit value for R is set up when the chip is programmed, and can be any random number except 0 (an LFSR filled with 0s will produce a never-ending stream of 0s). R is transformed by XORing bits 0, 2, 3, and 5 together, and shifting all 160 bits right 1 bit using the XOR result as the input bit to b₁₅₉. The process is shown in FIG. 347 below.

Care should be taken when updating R since it lives in Flash. Program code must assume power could be removed at any time.

The Read command is implemented with the following steps:

-   -   Wait for MinTicksRemaining to become 0     -   Loop through all of Flash, reading each word (will trigger         checks)     -   Accept n     -   Accept t     -   Restrict n to N     -   Restrict t to T     -   Accept R_(E)     -   Advance R_(L)     -   Output R_(L)     -   Output M_(Lt)     -   Sig←S_(Kn)[R_(E)|R_(L)|C₁|M_(Lt)]# calculation must take         constant time     -   MinTicksRemaining←MinTicks     -   Output Sig     -   Wait for MinTicksRemaining to become 0         15.1.9 Set Permissions     -   Input: n, p, R_(E), P_(E), SIG_(E)=[1 byte, 1 byte, 20 bytes, 4         bytes, 20 bytes]     -   Output: P_(p)     -   Changes: P_(p), R_(L)

The SetPermissions command is used to securely update the contents of P_(p) (containing QA Chip permissions). The WriteAuth command only attempts to replace P_(p) if the new value is signed combined with our local R.

It is only possible to sign messages by knowing K_(n). This can be achieved by a call to the SignP command (because only a ChipS can know K_(n)). It means that without a chip that can be used to produce the required signature, a write of any value to P_(p) is not possible.

The process is very similar to Test, except that if the validation succeeds, the P_(E) input parameter is additionally ORed with the current value for P_(p). Note that this is an OR, and not a replace. Since the SetParms command only sets bits in P_(p), the effect is to allow the permission bits corresponding to M[n] to progress from RW to either MSR, NMSR, or RO.

The SetPermissions command is implemented with the following steps:

-   -   Wait for MinTicksRemaining to become 0     -   Loop through all of Flash, reading each word (will trigger         checks)     -   Accept n     -   Restrict n to N     -   Accept p     -   Restrict p to T+N     -   Accept R_(E)     -   Accept P_(E)     -   SIG_(L)←S_(Kn)[R_(L)|R_(E)|P_(E)|C₂]# calculation must take         constant time     -   Accept SIG_(E)     -   If (SIG_(E)=SIG_(L))     -   Update R_(L)     -   P_(P)←P_(P)         P_(E)     -   EndIf     -   Output P_(P)# success or failure will be determined by receiver     -   MinTicksRemaining ← MinTicks         15.1.10 ReplaceKey     -   Input: n, R_(E), V, SIG_(E)=[1 byte, 20 bytes, 20 bytes, 20         bytes]     -   Output: Boolean (0x76=failure, 0x89=success)     -   Changes: K_(n), M_(L), R_(L)

The ReplaceKey command is used to replace the specified key in the QA Chip flash memory. However K_(n) can only be replaced if the previous value is known. A return byte of 0x89 is produced if the key was successfully updated, while 0x76 is returned for failure.

A ReplaceKey command consists of the WR_(A) command opcode followed by 0x89, 0x76, and then the appropriate parameters. Note that the new key is not sent in the clear, it is sent encrypted with the signature of R_(L), R_(E) and C₃ (signed with the old key). The first two input parameters must be verified by generating a signature using the old key.

The ReplaceKey command is implemented with the following steps:

Loop through all of Flash, reading each word (will trigger checks) Accept n Restrict n to N Accept R_(E) # session key from ChipF Accept V # encrypted key SIG_(L)

S_(Kn)[R_(E)|V|C₃] # calculation must take constant time Accept SIG_(E) If (SIG_(L) = SIG_(E2)) # comparison must take constant time SIG_(L)

S_(Kn)[R_(L)|R_(E)|C₃] # calculation must take constant time Advance R_(L) K_(E)

SIG_(L) ⊕ V K_(n)

K_(E)  # involves storing (K_(E) ⊕ R_(K)) and (

K_(E) ⊕ R_(K)) Output 0x89 # success Else Output 0x76 # failure EndIf 15.1.11 SignM

-   -   Input: n, R_(X), R_(E), M_(E), SIG_(E), M_(desired)=[1 byte, 20         bytes, 20 bytes, 64 bytes,32 bytes]     -   Output: R_(L), M_(new), S_(Kn)[R_(E)|R_(l)|C₁|M_(new)]=[20         bytes, 64 bytes, 20 bytes]     -   Changes:R_(L)

Note: The SignM command is only implemented in ChipS, and not in all QA Chips. The SignM command is used to produce a valid signed M for use in an authenticated write transaction. Only an QA Chip programmed with correct value of K_(n) can respond correctly to the SignM request. The output bytestream from the SignM command can be fed as the input bytestream to the WriteA command on a different QA Chip.

The input bytestream consists of the SMR opcode followed by 1 byte containing the key number to use for generating the signature, 20 bytes of Rx (representing the number passed in as R to ChipU's READ command, i.e. typically 0), the output from the READ command (namely R_(E), M_(E), and SIG_(E)), and finally the desired M to write to ChipU.

The SignM command only succeeds when SIG_(E)=S_(K)[R_(X)|R_(E)|C₁|M_(E)], indicating that the request was generated from a chip that knows K. This generation and comparison must take the same amount of time regardless of whether the input parameters are correct or not. If the times are not the same, an attacker can gain information about which bits of the supplied signature are incorrect. If the signatures match, then R_(L) is updated to be the next random number in the sequence.

Since the SignM function generates signatures, the function must wait for the MinTicksRemaining register to reach 0 before processing takes place.

Once all the inputs have been verified, a new memory vector is produced by applying a specially stored P value (eg word 1 of M₀) and M_(desired) against M_(E). Effectively, it is performing a regular Write, but with separate P against someone else's M. The M_(new) is signed with an updated R_(L) (and the passed in R_(E)), and all three values are output (the random number R_(L), M_(new), and the signature). The time taken to generate this signature must be the same regardless of the inputs.

Typically, the SignM command will be acting as a form of consumable command, so that a given ChipS can only generate a given number of signatures. The actual number can conveniently be stored in M₀ (eg word 0 of M₀) with ReadOnly permissions. Of course another chip could perform an Authorised write to update the number (using another ChipS) should it be desired.

The SignM command is implemented with the following steps:

Wait for MinTicksRemaining to become 0 Loop through all of Flash, reading each word (will trigger checks) Accept n Restrict n to N Accept R_(X) # don't care what this number is Accept R_(E) Accept M_(E) SIG_(L)

S_(Kn)[R_(X)|R_(E)|C₁|M_(E)] # calculation must take constant time Accept SIG_(E) Accept M_(desired) If ((SIG_(E) ≠ SIG_(L)) OR (M_(L)[0] = 0)) # fail if bad signature or if allowed sigs = 0 Output appropriate number of 0 # report failure Done EndIf Update R_(L) # Create the new version of M in ram from W and Permissions # This is the same as the core process of Write function # except that we don't write the results back to M DecEncountered

0 EqEncountered

0 Permissions = M_(L)[1] # assuming M₀ contains appropriate permissions For n

msw to lsw #(word 15 to 0) AM

Permissions[n] LT

(M_(desired)[n] < M_(E)[n]) # comparison is unsigned EQ

(M_(desired)[n] = M_(E)[n]) WE 

 (AM = RW) 

 ((AM = MSR) 

 LT) 

((AM = NMSR) 

(DecEncountered

LT)) DecEncountered

((AM = MSR)

LT)

((AM = NMSR)

DecEncountered)

((AM = NMSR)

EqEncountered

LT) EqEncountered 

 ((AM = MSR) 

 EQ) 

 ((AM = NMSR) 

EqEncountered

EQ) If (

WE)

(M_(E)[n] ≠ M_(desired)[n]) Output appropriate number of 0 # report failure EndIf EndFor # At this point, M_(desired) is correct Output R_(L) Output M_(desired) # M_(desired) is now effectively M_(new) Sig

S_(Kn)[R_(E)|R_(L)|C₁|M_(desired)] # calculation must take constant time MinTicksRemaining

MinTicks Decrement M_(L)[0] # reduce the number of allowable signatures by 1 Output Sig 15.1.12 SignP

-   -   Input: n,R_(E),P_(desired)=[1 byte, 20 bytes, 4 bytes]     -   Output: R_(L), S_(Kn)[R_(E)|R_(L)|P_(desired)|C₂]=[20 bytes, 20         bytes]     -   Changes: R_(L)

Note: The SignP command is only implemented in ChipS, and not in all QA Chips.

The SignP command is used to produce a valid signed P for use in a SetPermissions transaction.

Only an QA Chip programmed with correct value of K_(n) can respond correctly to the SignP request. The output bytestream from the SignP command can be fed as the input bytestream to the SetPermissions command on a different QA Chip.

The input bytestream consists of the SMP opcode followed by 1 byte containing the key number to use for generating the signature, 20 bytes of R_(E) (representing the number obtained from ChipU's RND command, and finally the desired P to write to ChipU.

Since the SignP function generates signatures, the function must wait for the MinTicksRemaining register to reach 0 before processing takes place.

Once all the inputs have been verified, the P_(desired) is signed with an updated R_(L) (and the passed in R_(E)), and both values are output (the random number R_(L) and the signature). The time taken to generate this signature must be the same regardless of the inputs.

Typically, the SignP command will be acting as a form of consumable command, so that a given ChipS can only generate a given number of signatures. The actual number can conveniently be stored in M₀ (eg word 0 of M₀) with ReadOnly permissions. Of course another chip could perform an Authorised write to update the number (using another ChipS) should it be desired.

The SignM command is implemented with the following steps:

Wait for MinTicksRemaining to become 0 Loop through all of Flash, reading each word (will trigger checks) Accept n Restrict n to N Accept R_(E) Accept P_(desired) If (M_(L)[0] = 0) # fail if allowed sigs = 0 Output appropriate number of 0 # report failure Done EndIf Update R_(L) Output R_(L) Sig

S_(Kn)[R_(E)|R_(L)|P_(desired)|C₂]  # calculation must take constant time MinTicksRemaining

MinTicks Decrement M_(L)[0] # reduce the number of allowable signatures by 1 Output Sig 15.1.13 Test

-   -   Input: n, R_(E), M_(E), SIG_(E)=[1 byte, 20 bytes, 64 bytes, 20         bytes]     -   Output: Boolean (0x76=failure, 0x89=success)     -   Changes: R_(L)

The Test command is used to authenticate a read of an M from a non-trusted QA Chip.

The Test command consists of the TST command opcode followed by input parameters: n, R_(E), M_(E), and SIG_(E). The byte order is least significant byte to most significant byte for each command component. All but the first input parameter bytes are obtained as the output bytes from a Read command to a non-trusted QA Chip. The entire data does not have to be stored by the client. Instead, the bytes can be passed directly to the trusted QA Chip's Test command, and only M should be kept from the Read.

Calls to Test must wait for the MinTicksRemaining register to reach 0.

S_(Kn)[R_(L)|R_(E)|C₁|M_(E)] is then calculated, and compared against the input signature SIG_(E). If they are different, R_(L) is not changed, and 0x76 is returned to indicate failure. If they are the same, then R_(L) is updated to be the next random number in the sequence and 0x89 is returned to indicate success. Updating R_(L) only after success forces the caller to use a new random number (via the Random command) each time a successful authentication is performed.

The calculation of S_(Kn)[R_(L)|R_(E)|C₁|M_(E)] and the comparison against SIG_(E) must take identical time so that the time to evaluate the comparison in the TST function is always the same. Thus no attacker can compare execution times or number of bits processed before an output is given.

The Test command is implemented with the following steps:

Wait for MinTicksRemaining to become 0 Loop through all of Flash, reading each word (will trigger checks) Accept n Restrict n to N Accept R_(E) Accept M_(E) SIG_(L)

S_(Kn)[R_(L)|R_(E)|C₁|M_(E)] # calculation must take constant time Accept SIG_(E) If (SIG_(E) = SIG_(L)) Update R_(L) Output 0x89 # success Else Output 0x76 # report failure EndIf MinTicksRemaining

MinTicks 15.1.14 Write

-   -   Input: t, M_(new), SIG_(E)=[1 byte, 64 bytes, 20 bytes]     -   Output: Boolean (0x76=failure, 0x89=success)     -   Changes: M_(t)

The Write command is used to update M_(t) according to the permissions in P_(t). The WR command by itself is not secure, since a clone QA Chip may simply return success every time. Therefore a Write command should be followed by an authenticated read of M_(t) (e.g. via a Read command) to ensure that the change was actually made.

The Write command is called by passing the WR command opcode followed by which M to be updated, the new data to be written to M, and a digital signature of M. The data is sent least significant byte to most significant byte.

The ability to write to a specific 32-bit word within M_(t) is governed by the corresponding Permissions bits as stored in P_(t). P_(t) can be set using the SetPermissions command.

The fact that M_(t) is Flash memory must be taken into account when writing the new value to M. It is possible for an attacker to remove power at any time. In addition, only the changes to M should be stored for maximum utilization. In addition, the longevity of M will need to be taken into account. This may result in the location of M being updated.

The signature is not keyed, since it must be generated by the consumable user.

The Write command is implemented with the following steps:

Loop through all of Flash, reading each word (will trigger checks) Accept t Restrict t to T Accept M_(E) # new M Accept SIG_(E) SIG_(L) = Generate SHA1[M_(E)] If (SIG_(L) = SIG_(E)) output 0x76 # failure due to invalid signature exit EndIf DecEncountered

0 EqEncountered

0 For i

msw to lsw #(word 15 to 0) P

P_(t)[i] LT

(M_(E)[i] < M_(t)[i]) # comparison is unsigned EQ

(M_(E)[i] = M_(t)[i])

(DecEncountered

LT)) DecEncountered

((P = MSR)

LT)

((P = NMSR)

DecEncountered)

((P = NMSR)

EqEncountered

LT) EqEncountered

((P = MSR)

EQ)

((P = NMSR)

EqEncountered

EQ) If (

WE)

(M_(E)[i] ≠ M_(t)[i]) output 0x76 # failure due to wanting a change but not allowed it EndIf EndFor # At this point, M_(E) (desired) is correct to be written to the flash M_(t)

M_(E) # update flash output 0x89 # success 15.1.15 WriteAuth

-   -   Input: n, R_(E), M_(E), SIG_(E)=[1 byte, 20 bytes, 64 bytes, 20         bytes]     -   Output: Boolean (0x76=failure, 0x89=success)     -   Changes: M₀, R_(L)

The WriteAuth command is used to securely replace the entire contents of M₀ (containing QA Chip application specific data) according to the P_(T+n). The WriteAuth command only attempts to replace M₀ if the new value is signed combined with our local R.

It is only possible to sign messages by knowing K_(n). This can be achieved by a call to the SignM command (because only a ChipS can know K_(n)). It means that without a chip that can be used to produce the required signature, a write of any value to M₀ is not possible.

The process is very similar to Write, except that if the validation succeeds, the M_(E) input parameter is processed against M₀ using permissions P_(T+n).

The WriteAuth command is implemented with the following steps:

Wait for MinTicksRemaining to become 0 Loop through all of Flash, reading each word (will trigger checks) Accept n Restrict n to N Accept R_(E) Accept M_(E) SIG_(L)

S_(Kn)[R_(L)|R_(E)|C₁|M_(E)] # calculation must take constant time Accept SIG_(E) If (SIG_(E) = SIG_(L)) Update R_(L) DecEncountered

0 EqEncountered

0 For i

msw to lsw #(word 15 to 0) P

P_(T+n)[i] LT

(M_(E)[i] < M₀[i]) # comparison is unsigned EQ

(M_(E)[i] = M₀[i]) WE 

 (P = RW) 

 ((P = MSR) 

 LT) 

((P = NMSR) 

(DecEncountered

LT)) DecEncountered

((P = MSR)

LT)

((P = NMSR)

DecEncountered)

((P = NMSR)

EqEncountered

LT) EqEncountered 

 ((P = MSR) 

 EQ) 

((P = NMSR) 

EqEncountered

EQ) If ((

WE)

(M_(E)[i] ≠ M₀[i])) output 0x76 # failure due to wanting a change but not allowed it EndIf EndFor # At this point, M_(E) (desired) is correct to be written to the flash M₀

M_(E) # update flash output 0x89 # success EndIf MinTicksRemaining

MinTicks 16 Manufacture

This chapter makes some general comments about the manufacture and implementation of authentication chips. While the comments presented here are general, see [84] for a detailed description of an implementation of an authentication chip.

The authentication chip algorithms do not constitute a strong encryption device. The net effect is that they can be safely manufactured in any country (including the USA) and exported to anywhere in the world.

The circuitry of the authentication chip must be resistant to physical attack. A summary of manufacturing implementation guidelines is presented, followed by specification of the chip's physical defenses (ordered by attack).

Note that manufacturing comments are in addition to any legal protection undertaken, such as patents, copyright, and license agreements (for example, penalties if caught reverse engineering the authentication chip).

16.1 Guidelines For Manufacturing

The following are general guidelines for implementation of an authentication chip in terms of manufacture (see [84] for a detailed description of an authentication chip). No special security is required during the manufacturing process.

-   -   Standard process     -   Minimum size (if possible)     -   Clock Filter     -   Noise Generator     -   Tamper Prevention and Detection circuitry     -   Protected memory with tamper detection     -   Boot circuitry for loading program code     -   Special implementation of FETs for key data paths     -   Data connections in polysilicon layers where possible     -   OverUnderPower Detection Unit     -   No test circuitry     -   Transparent epoxy packaging

Finally, as a general note to manufacturers of Systems, the data line to the System authentication chip and the data line to the Consumable authentication chip must not be the same line. See Section 16.2.3 on page 736.

16.1.1 Standard Process

The authentication chip should be implemented with a standard manufacturing process (such as Flash). This is necessary to:

-   -   allow a great range of manufacturing location options     -   take advantage of well-defined and well-behaved technology     -   reduce cost

Note that the standard process still allows physical protection mechanisms.

16.1.2 Minimum Size

The authentication chip must have a low manufacturing cost in order to be included as the authentication mechanism for low cost consumables. It is therefore desirable to keep the chip size as low as reasonably possible.

Each authentication chip requires 962 bits of non-volatile memory. In addition, the storage required for optimized HMAC-SHA1 is 1024 bits. The remainder of the chip (state machine, processor, CPU or whatever is chosen to implement Protocol C1) must be kept to a minimum in order that the number of transistors is minimized and thus the cost per chip is minimized. The circuit areas that process the secret key information or could reveal information about the key should also be minimized (see Section 16.1.8 on page 734 for special data paths).

16.1.3 Clock Filter

The authentication chip circuitry is designed to operate within a specific clock speed range. Since the user directly supplies the clock signal, it is possible for an attacker to attempt to introduce race-conditions in the circuitry at specific times during processing. An example of this is where a high clock speed (higher than the circuitry is designed for) may prevent an XOR from working properly, and of the two inputs, the first may always be returned. These styles of transient fault attacks can be very efficient at recovering secret key information, and have been documented in [5] and [1]. The lesson to be learned from this is that the input clock signal cannot be trusted.

Since the input clock signal cannot be trusted, it must be limited to operate up to a maximum frequency. This can be achieved a number of ways.

One way to filter the clock signal is to use an edge detect unit passing the edge on to a delay, which in turn enables the input clock signal to pass through.

FIG. 348 shows clock signal flow within the Clock Filter.

The delay should be set so that the maximum clock speed is a particular frequency (e.g. about 4 MHz). Note that this delay is not programmable—it is fixed.

The filtered clock signal would be further divided internally as required.

16.1.4 Noise Generator

Each authentication chip should contain a noise generator that generates continuous circuit noise. The noise will interfere with other electromagnetic emissions from the chip's regular activities and add noise to the I_(dd) signal. Placement of the noise generator is not an issue on an authentication chip due to the length of the emission wavelengths.

The noise generator is used to generate electronic noise, multiple state changes each clock cycle, and as a source of pseudo-random bits for the Tamper Prevention and Detection circuitry (see Section 16.1.5 on page 731).

A simple implementation of a noise generator is a 64-bit maximal period LFSR seeded with a non-zero number. The clock used for the noise generator should be running at the maximum clock rate for the chip in order to generate as much noise as possible.

16.1.5 Tamper Prevention and Detection Circuitry

A set of circuits is required to test for and prevent physical attacks on the authentication chip. However what is actually detected as an attack may not be an intentional physical attack. It is therefore important to distinguish between these two types of attacks in an authentication chip:

-   -   where you can be certain that a physical attack has occurred.     -   where you cannot be certain that a physical attack has occurred.

The two types of detection differ in what is performed as a result of the detection. In the first case, where the circuitry can be certain that a true physical attack has occurred, erasure of Flash memory key information is a sensible action. In the second case, where the circuitry cannot be sure if an attack has occurred, there is still certainly something wrong. Action must be taken, but the action should not be the erasure of secret key information. A suitable action to take in the second case is a chip RESET. If what was detected was an attack that has permanently damaged the chip, the same conditions will occur next time and the chip will RESET again. If, on the other hand, what was detected was part of the normal operating environment of the chip, a RESET will not harm the key.

A good example of an event that circuitry cannot have knowledge about, is a power glitch. The glitch may be an intentional attack, attempting to reveal information about the key. It may, however, be the result of a faulty connection, or simply the start of a power-down sequence. It is therefore best to only RESET the chip, and not erase the key. If the chip was powering down, nothing is lost. If the System is faulty, repeated RESETs will cause the consumer to get the System repaired. In both cases the consumable is still intact.

A good example of an event that circuitry can have knowledge about, is the cutting of a data line within the chip. If this attack is somehow detected, it could only be a result of a faulty chip (manufacturing defect) or an attack. In either case, the erasure of the secret information is a sensible step to take.

Consequently each authentication chip should have 2 Tamper Detection Lines—one for definite attacks, and one for possible attacks. Connected to these Tamper Detection Lines would be a number of Tamper Detection test units, each testing for different forms of tampering. In addition, we want to ensure that the Tamper Detection Lines and Circuits themselves cannot also be tampered with.

At one end of the Tamper Detection Line is a source of pseudo-random bits (clocking at high speed compared to the general operating circuitry). The Noise Generator circuit described above is an adequate source. The generated bits pass through two different paths—one carries the original data, and the other carries the inverse of the data. The wires carrying these bits are in the layer above the general chip circuitry (for example, the memory, the key manipulation circuitry etc.). The wires must also cover the random bit generator. The bits are recombined at a number of places via an XOR gate. If the bits are different (they should be), a 1 is output, and used by the particular unit (for example, each output bit from a memory read should be ANDed with this bit value). The lines finally come together at the Flash memory Erase circuit, where a complete erasure is triggered by a 0 from the XOR. Attached to the line is a number of triggers, each detecting a physical attack on the chip. Each trigger has an oversize nMOS transistor attached to GND. The Tamper Detection Line physically goes through this nMOS transistor. If the test fails, the trigger causes the Tamper Detect Line to become 0. The XOR test will therefore fail on either this clock cycle or the next one (on average), thus RESETing or erasing the chip.

FIG. 349 illustrates the basic principle of a Tamper Detection Line in terms of tests and the XOR connected to either the Erase or RESET circuitry.

The Tamper Detection Line must go through the drain of an output transistor for each test, as illustrated by FIG. 350:

It is not possible to break the Tamper Detect Line since this would stop the flow of 1 s and 0s from the random source. The XOR tests would therefore fail. As the Tamper Detect Line physically passes through each test, it is not possible to eliminate any particular test without breaking the Tamper Detect Line.

It is important that the XORs take values from a variety of places along the Tamper Detect Lines in order to reduce the chances of an attack. FIG. 351 illustrates the taking of multiple XORs from the Tamper Detect Line to be used in the different parts of the chip. Each of these XORs can be considered to be generating a ChipOK bit that can be used within each unit or sub-unit.

A sample usage would be to have an OK bit in each unit that is ANDed with a given ChipOK bit each cycle. The OK bit is loaded with 1 on a RESET. If OK is 0, that unit will fail until the next RESET. If the Tamper Detect Line is functioning correctly, the chip will either RESET or erase all key information. If the RESET or erase circuitry has been destroyed, then this unit will not function, thus thwarting an attacker.

The destination of the RESET and Erase line and associated circuitry is very context sensitive. It needs to be protected in much the same way as the individual tamper tests. There is no point generating a RESET pulse if the attacker can simply cut the wire leading to the RESET circuitry. The actual implementation will depend very much on what is to be cleared at RESET, and how those items are cleared.

Finally, FIG. 352 shows how the Tamper Lines cover the noise generator circuitry of the chip. The generator and NOT gate are on one level, while the Tamper Detect Lines run on a level above the generator.

16.1.6 Protected Memory with Tamper Detection

It is not enough to simply store secret information or program code in Flash memory. The Flash memory and RAM must be protected from an attacker who would attempt to modify (or set) a particular bit of program code or key information. The mechanism used must conform to being used in the Tamper Detection Circuitry (described above).

The first part of the solution is to ensure that the Tamper Detection Line passes directly above each Flash or RAM bit. This ensures that an attacker cannot probe the contents of Flash or RAM. A breach of the covering wire is a break in the Tamper Detection Line. The breach causes the Erase signal to be set, thus deleting any contents of the memory. The high frequency noise on the Tamper Detection Line also obscures passive observation.

The second part of the solution for Flash is to use multi-level data storage, but only to use a subset of those multiple levels for valid bit representations. Normally, when multi-level Flash storage is used, a single floating gate holds more than one bit. For example, a 4-voltage-state transistor can represent two bits. Assuming a minimum and maximum voltage representing 00 and 11 respectively, the two middle voltages represent 01 and 10. In the authentication chip, we can use the two middle voltages to represent a single bit, and consider the two extremes to be invalid states. If an attacker attempts to force the state of a bit one way or the other by closing or cutting the gate's circuit, an invalid voltage (and hence invalid state) results.

The second part of the solution for RAM is to use a parity bit. The data part of the register can be checked against the parity bit (which will not match after an attack).

The bits coming from Flash and RAM can therefore be validated by a number of test units (one per bit) connected to the common Tamper Detection Line. The Tamper Detection circuitry would be the first circuitry the data passes through (thus stopping an attacker from cutting the data lines).

While the multi-level Flash protection is enough for non-secret information, such as program code, R, and MinTicks, it is not sufficient for protecting K₁ and K₂. If an attacker adds electrons to a gate (see Section 5.7.2.15 on page 656) representing a single bit of K₁, and the chip boots up yet doesn't activate the Tamper Detection Line, the key bit must have been a 0. If it does activate the Tamper Detection Line, it must have been a 1. For this reason, all other non-volatile memory can activate the Tamper Detection Line, but K₁ and K₂ must not. Consequently Checksum is used to check for tampering of K₁ and K₂. A signature of the expanded form of K₁ and K₂ (i.e. 320 bits instead of 160 bits for each of K₁ and K₂) is produced, and the result compared against the Checksum. Any non-match causes a clear of all key information.

16.1.7 Boot Circuitry for Loading Program Code

Program code should be kept in multi-level Flash instead of ROM, since ROM is subject to being altered in a non-testable way. A boot mechanism is therefore required to load the program code into Flash memory (Flash memory is in an indeterminate state after manufacture).

The boot circuitry must not be in ROM—a small state-machine would suffice. Otherwise the boot code could be modified in an undetectable way.

The boot circuitry must erase all Flash memory, check to ensure the erasure worked, and then load the program code. Flash memory must be erased before loading the program code. Otherwise an attacker could put the chip into the boot state, and then load program code that simply extracted the existing keys. The state machine must also check to ensure that all Flash memory has been cleared (to ensure that an attacker has not cut the Erase line) before loading the new program code.

The loading of program code must be undertaken by the secure Programming Station before secret information (such as keys) can be loaded. This step must be undertaken as the first part of the programming process.

16.1.8 Special Implementation of FETs for Key Data Paths

The normal situation for FET implementation for the case of a CMOS Inverter (which involves a pMOS transistor combined with an nMOS transistor) as shown in FIG. 353:

During the transition, there is a small period of time where both the nMOS transistor and the pMOS transistor have an intermediate resistance. The resultant power-ground short circuit causes a temporary increase in the current, and in fact accounts for the majority of current consumed by a CMOS device. A small amount of infrared light is emitted during the short circuit, and can be viewed through the silicon substrate (silicon is transparent to infrared light). A small amount of light is also emitted during the charging and discharging of the transistor gate capacitance and transmission line capacitance.

For circuitry that manipulates secret key information, such information must be kept hidden. An alternative non-flashing CMOS implementation should therefore be used for all data paths that manipulate the key or a partially calculated value that is based on the key.

The use of two non-overlapping clocks φ1 and φ2 can provide a non-flashing mechanism. φ1 is connected to a second gate of all nMOS transistors, and φ2 is connected to a second gate of all pMOS transistors. The transition can only take place in combination with the clock. Since φ1 and φ2 are non-overlapping, the pMOS and nMOS transistors will not have a simultaneous intermediate resistance. The setup is shown in FIG. 354:

Finally, regular CMOS inverters can be positioned near critical non-Flashing CMOS components. These inverters should take their input signal from the Tamper Detection Line above. Since the Tamper Detection Line operates multiple times faster than the regular operating circuitry, the net effect will be a high rate of light-bursts next to each non-Flashing CMOS component. Since a bright light overwhelms observation of a nearby faint light, an observer will not be able to detect what switching operations are occurring in the chip proper. These regular CMOS inverters will also effectively increase the amount of circuit noise, reducing the SNR and obscuring useful EMI.

There are a number of side effects due to the use of non-Flashing CMOS:

-   -   The effective speed of the chip is reduced by twice the rise         time of the clock per clock cycle. This is not a problem for an         authentication chip.     -   The amount of current drawn by the non-Flashing CMOS is reduced         (since the short circuits do not occur). However, this is offset         by the use of regular CMOS inverters.     -   Routing of the clocks increases chip area, especially since         multiple versions of φ1 and φ2 are required to cater for         different levels of propagation. The estimation of chip area is         double that of a regular implementation.     -   Design of the non-Flashing areas of the authentication chip are         slightly more complex than to do the same with a with a regular         CMOS design. In particular, standard cell components cannot be         used, making these areas full custom. This is not a problem for         something as small as an authentication chip, particularly when         the entire chip does not have to be protected in this manner.         16.1.9 Connections in Polysilicon Layers Where Possible

Wherever possible, the connections along which the key or secret data flows, should be made in the polysilicon layers. Where necessary, they can be in metal 1, but must never be in the top metal layer (containing the Tamper Detection Lines).

16.1.10 OverUnderPower Detection Unit

Each authentication chip requires an OverUnderPower Detection Unit to prevent Power Supply Attacks. An OverUnderPower Detection Unit detects power glitches and tests the power level against a Voltage Reference to ensure it is within a certain tolerance. The Unit contains a single Voltage Reference and two comparators. The OverUnderPower Detection Unit would be connected into the RESET Tamper Detection Line, thus causing a RESET when triggered.

A side effect of the OverUnderPower Detection Unit is that as the voltage drops during a power-down, a RESET is triggered, thus erasing any work registers.

16.1.11 No Test Circuitry

Test hardware on an authentication chip could very easily introduce vulnerabilities. As a result, the authentication chip should not contain any BIST or scan paths.

The authentication chip must therefore be testable with external test vectors. This should be possible since the authentication chip is not complex.

16.1.12 Transparent Epoxy Packaging

The authentication chip needs to be packaged in transparent epoxy so it can be photo-imaged by the programming station to prevent Trojan horse attacks. The transparent packaging does not compromise the security of the authentication chip since an attacker can fairly easily remove a chip from its packaging. For more information see Section 16.2.20 on page 743 and [86].

16.2 Resistance To Physical Attacks

While this chapter only describes manufacture in general terms (since this document does not cover a specific implementation of a Protocol C1 authentication chip), we can still make some observations about such a chip's resistance to physical attack. A description of the general form of each physical attack can be found in Section 5.7.2 on page 652.

16.2.1 Reading ROM

This attack depends on the key being stored in an addressable ROM. Since each authentication chip stores its authentication keys in internal Flash memory and not in an addressable ROM, this attack is irrelevant.

16.2.2 Reverse Engineering the Chip

Reverse engineering a chip is only useful when the security of authentication lies in the algorithm alone. However our authentication chips rely on a secret key, and not in the secrecy of the algorithm. Our authentication algorithm is, by contrast, public, and in any case, an attacker of a high volume consumable is assumed to have been able to obtain detailed plans of the internals of the chip.

In light of these factors, reverse engineering the chip itself, as opposed to the stored data, poses no threat.

16.2.3 Usurping the Authentication Process

There are several forms this attack can take, each with varying degrees of success. In all cases, it is assumed that a clone manufacturer will have access to both the System and the consumable designs.

An attacker may attempt to build a chip that tricks the System into returning a valid code instead of generating an authentication code. This attack is not possible for two reasons. The first reason is that System authentication chips and Consumable authentication chips, although physically identical, are programmed differently. In particular, the RD opcode and the RND opcode are the same, as are the WR and TST opcodes. A System authentication Chip cannot perform a RD command since every call is interpreted as a call to RND instead. The second reason this attack would fail is that separate serial data lines are provided from the System to the System and Consumable authentication chips. Consequently neither chip can see what is being transmitted to or received from the other.

If the attacker builds a clone chip that ignores WR commands (which decrement the consumable remaining), Protocol C1 ensures that the subsequent RD will detect that the WR did not occur. The System will therefore not go ahead with the use of the consumable, thus thwarting the attacker. The same is true if an attacker simulates loss of contact before authentication—since the authentication does not take place, the use of the consumable doesn't occur.

An attacker is therefore limited to modifying each System in order for clone consumables to be accepted (see Section 16.2.4 on page 737 for details of resistance this attack).

16.2.4 Modification of System

The simplest method of modification is to replace the System's authentication chip with one that simply reports success for each call to TST. This can be thwarted by System calling TST several times for each authentication, with the first few times providing false values, and expecting a fail from TST. The final call to TST would be expected to succeed. The number of false calls to TST could be determined by some part of the returned result from RD or from the system clock. Unfortunately an attacker could simply rewire System so that the new System clone authentication chip can monitor the returned result from the consumable chip or clock. The clone System authentication chip would only return success when that monitored value is presented to its TST function. Clone consumables could then return any value as the hash result for RD, as the clone System chip would declare that value valid. There is therefore no point for the System to call the System authentication chip multiple times, since a rewiring attack will only work for the System that has been rewired, and not for all Systems.

A similar form of attack on a System is a replacement of the System ROM. The ROM program code can be altered so that the Authentication never occurs. There is nothing that can be done about this, since the System remains in the hands of a consumer. Of course this would void any warranty, but the consumer may consider the alteration worthwhile if the clone consumable were extremely cheap and more readily available than the original item.

The System/consumable manufacturer must therefore determine how likely an attack of this nature is. Such a study must include given the pricing structure of Systems and Consumables, frequency of System service, advantage to the consumer of having a physical modification performed, and where consumers would go to get the modification performed.

The likelihood of physical alteration increases with the perceived artificiality of the consumable marketing scheme. It is one thing for a consumable to be protected against clone manufacturers. It is quite another for a consumable's market to be protected by a form of exclusive licensing arrangement that creates what is viewed by consumers as artificial markets. In the former case, owners are not so likely to go to the trouble of modifying their system to allow a clone manufacturer's goods. In the latter case, consumers are far more likely to modify their System. A case in point is DVD. Each DVD is marked with a region code, and will only play in a DVD player from that region. Thus a DVD from the USA will not play in an Australian player, and a DVD from Japan, Europe or Australia will not play in a USA DVD player. Given that certain DVD titles are not available in all regions, or because of quality differences, pricing differences or timing of releases, many consumers have had their DVD players modified to accept DVDs from any region. The modification is usually simple (it often involves soldering a single wire), voids the owner's warranty, and often costs the owner some money. But the interesting thing to note is that the change is not made so the consumer can use clone consumables—the consumer will still only buy real consumables, but from different regions. The modification is performed to remove what is viewed as an artificial barrier, placed on the consumer by the movie companies. In the same way, a System/Consumable scheme that is viewed as unfair will result in people making modifications to their Systems.

The limit case of modifying a system is for a clone manufacturer to provide a completely clone System which takes clone consumables. This may be simple competition or violation of patents. Either way, it is beyond the scope of the authentication chip and depends on the technology or service being cloned.

16.2.5 Direct Viewing of Chip Operation by Conventional Probing

In order to view the chip operation, the chip must be operating. However, the Tamper Prevention and Detection circuitry covers those sections of the chip that process or hold the key. It is not possible to view those sections through the Tamper Prevention lines.

An attacker cannot simply slice the chip past the Tamper Prevention layer, for this will break the Tamper Detection Lines and cause an erasure of all keys at power-up. Simply destroying the erasure circuitry is not sufficient, since the multiple ChipOK bits (now all 0) feeding into multiple units within the authentication chip will cause the chip's regular operating circuitry to stop functioning.

To set up the chip for an attack, then, requires the attacker to delete the Tamper Detection lines, stop the Erasure of Flash memory, and somehow rewire the components that relied on the ChipOK lines. Even if all this could be done, the act of slicing the chip to this level will most likely destroy the charge patterns in the non-volatile memory that holds the keys, making the process fruitless.

16.2.6 Direct Viewing of the Non-Volatile Memory

If the authentication chip were sliced so that the floating gates of the Flash memory were exposed, without discharging them, then the keys could probably be viewed directly using an STM or SKM. However, slicing the chip to this level without discharging the gates is probably impossible. Using wet etching, plasma etching, ion milling, or chemical mechanical polishing will almost certainly discharge the small charges present on the floating gates. This is true of regular Flash memory, but even more so of multi-level Flash memory.

16.2.7 Viewing the Light Bursts Caused by State Changes

All sections of circuitry that manipulate secret key information are implemented in the non-Flashing CMOS described above. This prevents the emission of the majority of light bursts. Regular CMOS inverters placed in close proximity to the non-Flashing CMOS will hide any faint emissions caused by capacitor charge and discharge. The inverters are connected to the Tamper Detection circuitry, so they change state many times (at the high clock rate) for each non-Flashing CMOS state change.

16.2.8 Viewing the Keys Using an SEPM

An SEPM attack can be simply thwarted by adding a metal layer to cover the circuitry. However an attacker could etch a hole in the layer, so this is not an appropriate defense.

The Tamper Detection circuitry described above will shield the signal as well as cause circuit noise. The noise will actually be a greater signal than the one that the attacker is looking for. If the attacker attempts to etch a hole in the noise circuitry covering the protected areas, the chip will not function, and the SEPM will not be able to read any data.

An SEPM attack is therefore fruitless.

16.2.9 Monitoring EMI

The Noise Generator described above will cause circuit noise. The noise will interfere with other electromagnetic emissions from the chip's regular activities and thus obscure any meaningful reading of internal data transfers.

16.2.10 Viewing I_(dd) Fluctuations

The solution against this kind of attack is to decrease the SNR in the I_(dd) signal. This is accomplished by increasing the amount of circuit noise and decreasing the amount of signal.

The Noise Generator circuit (which also acts as a defense against EMI attacks) will also cause enough state changes each cycle to obscure any meaningful information in the I_(dd) signal.

In addition, the special Non-Flashing CMOS implementation of the key-carrying data paths of the chip prevents current from flowing when state changes occur. This has the benefit of reducing the amount of signal.

16.2.11 Differential Fault Analysis

Differential fault bit errors are introduced in a non-targeted fashion by ionization, microwave radiation, and environmental stress. The most likely effect of an attack of this nature is a change in Flash memory (causing an invalid state) or RAM (bad parity). Invalid states and bad parity are detected by the Tamper Detection Circuitry, and cause an erasure of the key.

Since the Tamper Detection Lines cover the key manipulation circuitry, any error introduced in the key manipulation circuitry will be mirrored by an error in a Tamper Detection Line. If the Tamper Detection Line is affected, the chip will either continually RESET or simply erase the key upon a power-up, rendering the attack fruitless.

Rather than relying on a non-targeted attack and hoping that “just the right part of the chip is affected in just the right way”, an attacker is better off trying to introduce a targeted fault (such as overwrite attacks, gate destruction etc.). For information on these targeted fault attacks, see the relevant sections below.

16.2.12 Clock Glitch Attacks

The Clock Filter (described above) eliminates the possibility of clock glitch attacks.

16.2.13 Power Supply Attacks

The OverUnderPower Detection Unit (described above) eliminates the possibility of power supply attacks.

16.2.14 Overwriting ROM

Authentication chips store program code, keys and secret information in Flash memory, and not in ROM. This attack is therefore not possible.

16.2.15 Modifying EEPROM/Flash

Authentication chips store program code, keys and secret information in multi-level Flash memory. However the Flash memory is covered by two Tamper Prevention and Detection Lines. If either of these lines is broken (in the process of destroying a gate via a laser-cutter) the attack will be detected on power-up, and the chip will either RESET (continually) or erase the keys from Flash memory. This process is described in Section 16.1.6 on page 733.

Even if an attacker is able to somehow access the bits of Flash and destroy or short out the gate holding a particular bit, this will force the bit to have no charge or a full charge. These are both invalid states for the authentication chip's usage of the multi-level Flash memory (only the two middle states are valID). When that data value is transferred from Flash, detection circuitry will cause the Erasure Tamper Detection Line to be triggered—thereby erasing the remainder of Flash memory and RESETing the chip. This is true for program code, and non-secret information. As key data is read from multi-level flash memory, it is not imediately checked for validity (otherwise information about the key is given away). Instead, a specific key validation mechanism is used to protect the secret key information.

An attacker could theoretically etch off the upper levels of the chip, and deposit enough electrons to change the state of the multi-level Flash memory by ⅓. If the beam is high enough energy it might be possible to focus the electron beam through the Tamper Prevention and Detection Lines. As a result, the authentication chip must perform a validation of the keys before replying to the Random, Test or Random commands. The SHA-1 algorithm must be run on the keys, and the results compared against an internal checksum value. This gives an attacker a 1 in 2¹⁶⁰ chance of tricking the chip, which is the same chance as guessing either of the keys.

A Modify EEPROM/Flash attack is therefore fruitless.

16.2.16 Gate Destruction Attacks

Gate Destruction Attacks rely on the ability of an attacker to modify a single gate to cause the chip to reveal information during operation. However any circuitry that manipulates secret information is covered by one of the two Tamper Prevention and Detection lines. If either of these lines is broken (in the process of destroying a gate) the attack will be detected on power-up, and the chip will either RESET (continually) or erase the keys from Flash memory.

To launch this kind of attack, an attacker must first reverse-engineer the chip to determine which gate(s) should be targeted. Once the location of the target gates has been determined, the attacker must break the covering Tamper Detection line, stop the Erasure of Flash memory, and somehow rewire the components that rely on the ChipOK lines. Rewiring the circuitry cannot be done without slicing the chip, and even if it could be done, the act of slicing the chip to this level will most likely destroy the charge patterns in the non-volatile memory that holds the keys, making the process fruitless.

16.2.17 Overwrite Attack

An overwrite attack relies on being able to set individual bits of the key without knowing the previous value. It relies on probing the chip, as in the conventional probing attack and destroying gates as in the gate destruction attack. Both of these attacks (as explained in their respective sections), will not succeed due to the use of the Tamper Prevention and Detection Circuitry and ChipOK lines.

However, even if the attacker is able to somehow access the bits of Flash and destroy or short out the gate holding a particular bit, this will force the bit to have no charge or a full charge. These are both invalid states for the authentication chip's usage of the multi-level Flash memory (only the two middle states are valID). When that data value is transferred from Flash detection circuitry will cause the Erasure Tamper Detection Line to be triggered—thereby erasing the remainder of Flash memory and RESETing the chip. In the same way, a parity check on tampered values read from RAM will cause the Erasure Tamper Detection Line to be triggered.

An overwrite attack is therefore fruitless.

16.2.18 Memory Remanence Attack

Any working registers or RAM within the authentication chip may be holding part of the authentication keys when power is removed. The working registers and RAM would continue to hold the information for some time after the removal of power. If the chip were sliced so that the gates of the registers/RAM were exposed, without discharging them, then the data could probably be viewed directly using an STM.

The first defense can be found above, in the description of defense against power glitch attacks. When power is removed, all registers and RAM are cleared, just as the RESET condition causes a clearing of memory.

The chances then, are less for this attack to succeed than for a reading of the Flash memory. RAM charges (by nature) are more easily lost than Flash memory. The slicing of the chip to reveal the RAM will certainly cause the charges to be lost (if they haven't been lost simply due to the memory not being refreshed and the time taken to perform the slicing).

This attack is therefore fruitless.

16.2.19 Chip Theft Attack

There are distinct phases in the lifetime of an authentication chip. Chips can be stolen when at any of these stages:

-   -   After manufacture, but before programming of key     -   After programming of key, but before programming of state data     -   After programming of state data, but before insertion into the         consumable or system     -   After insertion into the system or consumable

A theft in between the chip manufacturer and programming station would only provide the clone manufacturer with blank chips. This merely compromises the sale of authentication chips, not anything authenticated by the authentication chips. Since the programming station is the only mechanism with consumable and system product keys, a clone manufacturer would not be able to program the chips with the correct key. Clone manufacturers would be able to program the blank chips for their own Systems and Consumables, but it would be difficult to place these items on the market without detection.

The second form of theft can only happen in a situation where an authentication chip passes through two or more distinct programming phases. This is possible, but unlikely. In any case, the worst situation is where no state data has been programmed, so all of M is read/write. If this were the case, an attacker could attempt to launch an adaptive chosen text attack on the chip. The HMAC-SHA1 algorithm is resistant to such attacks. For more information see Section 14.7 on page 699.

The third form of theft would have to take place in between the programming station and the installation factory. The authentication chips would already be programmed for use in a particular system or for use in a particular consumable. The only use these chips have to a thief is to place them into a clone System or clone Consumable. Clone systems are irrelevant—a cloned System would not even require an authentication chip. For clone Consumables, such a theft would limit the number of cloned products to the number of chips stolen. A single theft should not create a supply constant enough to provide clone manufacturers with a cost-effective business.

The final form of theft is where the System or Consumable itself is stolen. When the theft occurs at the manufacturer, physical security protocols must be enhanced. If the theft occurs anywhere else, it is a matter of concern only for the owner of the item and the police or insurance company. The security mechanisms that the authentication chip uses assume that the consumables and systems are in the hands of the public. Consequently, having them stolen makes no difference to the security of the keys.

16.2.20 Trojan Horse Attack

A Trojan horse attack involves an attacker inserting a fake authentication chip into the programming station and retrieving the same chip after it has been programmed with the secret key information. The difficulty of these two tasks depends on both logical and physical security, but is an expensive attack—the attacker has to manufacture a false authentication chip, and it will only be useful where the effort is worth the gain. For example, obtaining the secret key for a specific car's authentication chip is most likely not worth an attacker's efforts, while the key for a printer's ink cartridge may be very valuable.

The problem arises if the programming station is unable to tell a Trojan horse authentication chip from a real one—which is the problem of authenticating the authentication chip.

One solution to the authentication problem is for the manufacturer to have a programming station attached to the end of the production line. Chips passing the manufacture QA tests are programmed with the manufacturer's secret key information. The chip can therefore be verified by the C1 authentication protocol, and give information such as the expected batch number, serial number etc. The information can be verified and recorded, and the valid chip can then be reprogrammed with the System or Consumable key and state data. An attacker would have to substitute an authentication chip with a Trojan horse programmed with the manufacturer's secret key information and copied batch number data from the removed authentication chip. This is only possible if the manufacturer's secret key is compromised (the key is changed regularly and not known by a human) or if the physical security at the manufacturing plant is compromised at the end of the manufacturing chain.

Even if the solution described were to be undertaken, the possibility of a Trojan horse attack does not go away—it merely is removed to the manufacturer's physical location. A better solution requires no physical security at the manufacturing location.

The preferred solution then, is to use transparent epoxy on the chip's packaging and to image the chip before programming it. Once the chip has been mounted for programming it is in a known fixed orientation. It can therefore be high resolution photo-imaged and X-rayed from multiple directions, and the images compared against “signature”images. Any chip not matching the image signature is treated as a Trojan horse and rejected.

1 Refill of Ink in Printers—Printer Based Refill Device

1.1 Functional Purpose

The functional purpose of the printer based refill device is as follows:

-   -   To refill ink into printers by physically connecting the refill         device to the printer.     -   To ensure that the correct ink is used for the correct operation         of the printer (i.e. will not damage the printhead).     -   To ensure accurate measure of ink is transferred from the         refilling device to the printer during refills.     -   The refill device is controlled by the printer. Apart from the         QA Chip¹ the refill device has no other processing power.         ¹General Note: Througout this document, if secure refilling is         required then a physical QA Chip or any other virtual device         performing the QA Chip protocol can be used. Refer to [1].         1.2 Basic Components Of The Refill Device

FIG. 355 shows the components of the printer based refill device.

The printer based refill device will consist of following components:

-   -   An ink reservoir—which stores the ink. Each refill device will         allow ink reservoirs of various capacities. When the ink         reservoir empties out, it is replaced by another reservoir         containing more ink of the same type or different type or         refilled (for example through a refill station as described in         Section 2 and Section 3).     -   An ink output device—which dispenses ink to the printer being         refilled when physically connected to the printer.     -   A QA Chip and associated circuitry—which stores the amount of         ink in the reservoir along with the attributes of the ink in a         digital format.     -   The electrical connections to the QA Chip.     -   NB—No additional microprocessors are required to be present in         the refill device. Hence the refill device uses the processing         power of the printer to oversee the refilling process.     -   An ink transfer mechanism (optional) which controls the flow ink         from the refill device to the printer and is controlled by the         printer. Therefore the control connections for the ink transfer         mechanism will be connected to the printer.     -   Alternatively, the ink transfer mechanism could be in the         printer. Refer to Section 1.3.         1.3 Printer Description And Functions

Printers which will be refilled by these refilling devices must have the following components:

-   -   Microprocessor assembly which will control the refill procedure         as described Section 1.4. The microprocessor assembly will         access the QA Chip and ink transfer mechanism of the refill         device.     -   A QA Chip storing the ink amount remaining in the printer.     -   An optional ink transfer mechanism to control the flow of ink         from the refill device to the printer. This ink transfer         mechanism must be present in the printer if the refill device         doesn't have one of its own.         1.4 Operational Procedure

The operational procedure can be divided into two parts:

-   -   Refilling printers using the refill device.     -   Refilling of the ink reservoir in the refill device. See Section         2 and Section 3.         1.4.1 Refilling of Printers

FIG. 356 shows a printer being refilled by a printer based refill device. The ink transfer mechanism is located in the printer in this case. The ink transfer mechanism could be also located in the refill device as described in Section 1.2.

The following is a description for refilling of printers using the printer based refill device:

-   -   Ink output device from the refilling device is connected to the         printer.     -   The QA Chip electrical connection is connected to the printer.     -   The refill option is selected on the user interface of the         printer. The microprocessor assembly in the printer will then do         the following:

-   a. Read ink attributes (for example ink type, ink characteristics,     ink colour, ink manufacturer etc) stored in the QA Chip of the ink     reservoir unit. Refer to [1].

-   b. Compare the ink attributes as required by the printer for correct     operation. This may require reading of data from the QA Chip in the     printer.

-   c. Only if Step b is successful, then do the following:     -   i. Determine the amount of ink to be transferred by any or all         of the following means, ensuring that the reservoir has enough         ink for the transfer:     -   Fixed amount (e.g. based on a pre-programmed value or printer         model).     -   User-selectable amount.         -   ii. Decrement the amount of ink transferred from the QA Chip             in the refill station and increment the QA Chip in the             printer (which stores the amount of ink in the printer) with             corresponding ink amount.         -   iii. Command the ink transfer mechanism to release the ink             to the printer through the output device.             2 Home Use Refill Station             2.1 Functional Purpose

The functional purpose of the commercial refill station is as follows:

-   -   To refill ink into ink cartridges at home or in a small office.     -   Single ink cartridge is filled at a time.     -   To ensure that the correct ink present in the refill station is         transferred to the correct ink cartridge.     -   To ensure accurate measure of ink is transferred from the         refilling station to the ink cartridge during refills.     -   The refilling station provides the processing power required to         perform refills of ink cartridges.         2.2 Basic Components

FIG. 357 shows the components of a home refill station.

A home refill station will consist of one of the following ink refill units:

-   -   A single reservoir ink refill unit suitable for black ink (or         any other single colour).     -   A multi reservoir ink refill unit suitable for coloured ink for         example CMY (Cyan, Magenta, Yellow).         2.2.1 Ink Reservoir Unit

FIG. 358 shows the components of a three-ink reservoir unit.

The ink reservoir unit will consist of the following:

-   -   Multiple ink reservoirs or a single ink reservoir which stores         ink. Each refill station will allow ink reservoirs of various         capacities. When the ink reservoir empties out, it is replaced         by another reservoir containing more ink of the same or         different type or refilled (for example through a refill station         as described in Section 3).     -   A QA Chip and associated circuitry in each of the ink         reservoirs—which stores the amount of ink in the reservoir along         with the attributes of the ink.     -   The electrical connections to each of the QA Chips.         2.2.2 Ink Transfer Unit

The ink reservoir unit will consist of the following:

-   -   Ink output device from each ink reservoir.     -   The output ink transfer mechanism controls the flow ink from the         ink refill unit to the ink cartridge and is controlled by the         microprocessor assembly.     -   Final ink output devices to the cartridge interface assembly         2.2.3 Cartridge Interface Unit

This unit will provide the physical interface to the ink cartridges. Each ink cartridge interface unit will hold a single or multiple cartridges of particular physical dimension.

The cartridge interface unit can removed from the ink refill unit and replaced with another interface unit to cater for other physically different cartridges.

2.2.4 Microprocessor Assembly

The controls connections for the ink transfer mechanism and the electrical connections of the QA Chip are connected to the microprocessor assembly. The microprocessor assembly oversees and controls the refill process.

The microprocessor assembly will communicate with a user interface to accept commands and provide responses for various refill operations.

2.3 Ink Cartridge Description

Ink cartridges which will be refilled in a home refill station must have a QA Chip storing the following components:

-   -   Ink amount remaining.     -   Ink attributes (for example—ink type, ink characteristics, ink         colour, ink manufacturer).         2.4 Operational Procedure

The operational procedure can be divided into two parts:

-   -   Refilling of ink cartridges using the home refill station.     -   Refilling the ink reservoirs used in the refill station is         discussed in Section 3.         2.5 Refilling Of Ink Cartridges Using The Home Refill Station

FIG. 359 shows the refill of ink cartridges in a home refill station.

The following is a description for refilling of ink cartridges in the home refill station:

-   -   Load the ink cartridge into the cartridge interface unit of the         ink refill unit. This will connect the QA Chip of the ink         cartridge to the microprocessor assembly. It will also connect         the ink output device of the ink refill unit to the ink         cartridge.     -   The model number of the ink cartridge is read from the QA Chip         by the microprocessor assembly controlling the ink refill units.     -   The microprocessor assembly will determine whether the ink         refill unit is suitable for the ink cartridge model.     -   The refill option is selected on the microprocessor assembly         through the user interface. The microprocessor assembly will         then do the following:

-   a. Read ink attributes (for example ink type, ink characteristics,     ink colour, ink manufacturer etc) stored in the QA Chip of the ink     cartridge. Refer to [1].

-   b. Compare the read ink attributes to the ink attribute list in the     refill station. This may also require reading of the ink attributes     stored in the QA Chip of the ink reservoirs in the refill unit.

-   c. Only if Step b is successful, then do the following:     -   i. Determine the amount of ink to be transferred by any or all         of the following means, ensuring that the reservoir has enough         ink for the transfer:     -   Fixed amount (e.g. based on a pre-programmed value cartridge         model or reservoir type).     -   User-selectable amount.         -   ii. Check the ink reservoir in the ink refill unit has             adequate amount of ink to refill the ink cartridge         -   iii. Decrement the amount of ink transferred from the QA             Chip in the ink refill unit and increment the QA Chip in the             ink cartridge with corresponding ink amount.         -   iv. If incrementing of the QA Chip with ink amount is             successful then a command is sent to the ink transfer             mechanism to release the ink to the ink cartridge through             the output device.             3 Commercial Refill Station             3.1 Functional Purpose

The functional purpose of the commercial refill station is as follows:

-   -   To refill ink into ink cartridges that are taken to the refill         station for refilling.     -   Multiple ink cartridges of different models can be refilled.     -   To ensure that the correct ink present in the refill station is         transferred to the ink cartridge.     -   To ensure accurate measure of ink is transferred from the         refilling station to the ink cartridge during refills.     -   The refilling station provides all processing power required to         perform refills of ink cartridges.         3.2 Basic Components Of The Refill Station

FIG. 360 shows the components of a commercial refill station.

A commercial refill station will consist of multiple ink refill units controlled by a single microprocessor assembly. Each ink refill unit can refill a single ink cartridge at a time.

Each ink refill unit will consist of the following sub units:

-   -   Ink reservoir unit     -   Switch unit     -   Ink transfer unit     -   Multiple cartridge interface unit         3.2.1         Ink Reservoir Unit

FIG. 361 shows the components of a ink reservoir unit.

The ink reservoir unit will consist of the following:

-   -   Multiple ink reservoirs—which stores ink. Each refill device         will allow ink reservoirs of various capacities. When the ink         reservoir empties out, it is replaced by another reservoir         containing more ink of the same or different type or refilled.         Refer to Section 3.5.     -   A QA Chip and associated circuitry in each of the ink         reservoirs—which stores the amount of ink in the reservoir along         with the attributes of the ink in digital format.     -   The electrical connections of each of the QA Chips are connected         to the microprocessor assembly.         3.2.2 Switch Unit

This unit will switch the inks selected from different ink reservoirs to the ink transfer unit to be dispensed into ink cartridges.

The switch unit will prevent mixing of any residual ink left in dispensing devices after each ink cartridge is refilled.

3.2.3 Ink Transfer Unit

The ink reservoir unit will consist of the following:

-   -   Ink output device from each ink reservoir.     -   An output ink transfer mechanism which controls the flow ink         from the ink refill unit to the ink cartridge and is controlled         by the microprocessor assembly.     -   Final ink output devices to the multiple cartridge interface         assembly         3.2.4 Multiple Cartridge Interface Unit

This unit will provide the physical interface to the ink cartridges. Each ink cartridge interface will hold cartridges of different physical dimensions.

Each cartridge interface unit can provide an interface for about 20 physically different cartridges.

The cartridge interface unit can removed from the ink refill unit and replaced with another interface unit to cater for other physically different cartridges.

3.2.5 Microprocessor Assembly with a User Interface

The controls connections for the ink transfer mechanism and the electrical connections of the QA Chip are connected to the microprocessor assembly. The microprocessor assembly will oversee and control the refill process.

The microprocessor assembly will communicate with a user interface to accept commands and provide responses for various refill operations.

3.3 Ink Cartridge Description

Ink cartridges which will be refilled in a commercial refill station must have a QA Chip storing the following components:

-   -   Ink amount remaining.     -   Ink attributes (for example—ink type, ink characteristics, ink         colour, ink manufacturer).         3.4 Operational Procedure

The operational procedure can be divided into two parts:

-   -   Refilling of ink cartridges using the commercial refill station.     -   Refilling the ink reservoirs used in the refill station is         covered in Section 3.5.         3.4.1 Refilling Ink Cartridges Using the Commercial Refill         Station

FIG. 362 shows the refill of ink cartridges in a commercial refill station.

The following is a description for refilling of ink cartridges in the commercial refill station:

-   -   Load the ink cartridge into the multiple cartridge interface         unit of the ink refill unit. This will connect the QA Chip of         the ink cartridge to the microprocessor assembly. It will also         connect the ink output device of the ink refill unit to the ink         cartridge.     -   The model number of the ink cartridge automatically is read from         the QA Chip by the microprocessor assembly controlling the ink         refill units.     -   The microprocessor assembly will determine whether the ink         refill unit is suitable for the ink cartridge model.     -   The refill option is selected on the microprocessor assembly         through the user interface. The microprocessor assembly will         then do the following:

-   a. Read ink attributes (for example ink type, ink characteristics,     ink colour, ink manufacturer etc) stored in the QA Chip of the ink     cartridge. Refer to [1].

-   b. Compare the read ink attributes to the ink attribute list in the     refill station. This may also require reading of the ink attributes     stored in the QA Chip of the ink reservoirs in the refill unit.

-   c. Only if Step b is successful, then do the following:     -   i. Determine the amount of ink to be transferred by any or all         of the following means, ensuring that the reservoir has enough         ink for the transfer:     -   Fixed amount (e.g. based on a pre-programmed value, cartridge         model or reservoir type).     -   User-selectable amount.         -   ii. The microprocessor assembly will calculate the cost of             ink amount and interrogate the user for a payment             method—credit card or cash. If credit card option is             selected it will request a credit card number to be selected             and interface to a payment system to complete the             transaction before proceeding further.         -   iii. Decrement the amount of ink transferred from the QA             Chip in the ink refill unit and increment the QA Chip in the             ink cartridge with corresponding ink amount.         -   iv. If incrementing of the QA Chip with ink amount is             successful then a command is sent to the ink transfer             mechanism to release the ink to the ink cartridge through             the output device.             3.5 Refilling The Ink Reservoirs

The ink reservoirs of any ink refill device can be refilled recursively by the procedure described in Section 3.4.1, the only exception being the ink cartridge replaced by the ink reservoir.

3.6 Commercial Refill Station For A Production Environment

This refill station resembles a commercial refill station but fills multiple ink cartridges of the same type at the same time. This will serve as a filling station for new cartridges in a production environment.

Logical Interface Specification for Preferred Form of QA Chip

1 Introduction

This document defines the QA Chip Logical Interface, which provides authenticated manipulation of specific printer and consumable parameters. The interface is described in terms of data structures and the functions that manipulate them, together with examples of use. While the descriptions and examples are targetted towards the printer application, they are equally applicable in other domains.

2 Scope

The document describes the QA Chip Logical Interface as follows:

-   -   data structures and their uses (Section 5 to Section 9).     -   functions, including inputs, outputs, signature formats, and a         logical implementation sequence (Section 10 to Section 30).     -   typical functional sequences of printers and consumables, using         the functions and data structures of the interface (Section 31         to Section 32).

The QA Chip Logical Interface is a logical interface, and is therefore implementation independent.

Although this document does not cover implementation details on particular platforms, expected implementations include:

-   -   Software only     -   Off-the-shelf cryptographic hardware.     -   ASICs, such as SBR4320 [2] and SOPEC [3] for physical insertion         into printers and ink cartridges     -   Smart cards.         3 Nomenclature         3.1 Symbols

The following symbolic nomenclature is used throughout this document:

TABLE 246 Summary of symbolic nomenclature Symbol Description F[X] Function F, taking a single parameter X F[X, Y] Function F, taking two parameters, X and Y X | Y X concatenated with Y X

Y Bitwise X AND Y X

Y Bitwise X OR Y (inclusive-OR) X ⊕ Y Bitwise X XOR Y (exclusive-OR)

X Bitwise NOT X (complement) X

Y X is assigned the value Y X

{Y, Z} The domain of assignment inputs to X is Y and Z X = Y X is equal to Y X ≠ Y X is not equal to Y

X Decrement X by 1 (floor 0)

X Increment X by 1 (modulo register length) Erase X Erase Flash memory register X SetBits[X, Y] Set the bits of the Flash memory register X based on Y Z

ShiftRight[X, Y] Shift register X right one bit position, taking input bit from Y and placing the output bit in Z a.b Data field or member function ‘b’ in object a. 3.2 Pseudocode 3.2.1 Asynchronous

The following pseudocode:

-   -   var=expression         -   means the var signal or output is equal to the evaluation of             the expression.             3.2.2 Synchronous

The following pseudocode:

-   -   var←expression         means the var register is assigned the result of evaluating the         expression during this cycle.         3.2.3 Expression

Expressions are defined using the nomenclature in Table 246 above. Therefore:

-   -   var=(a=b)         is interpreted as the var signal is 1 if a is equal to b, and 0         otherwise.         4 Terms         4.1 QA Device and System

An instance of a QA Chip Logical Interface (on any platform) is a QA Device.

QA Devices cannot talk directly to each other. A System is a logical entity which has one or more QA Devices connected logically (or physically) to it, and calls the functions on the QA Devices. The system is considered secure and the program running on the system is considered to be trusted.

4.2 Types of QA Devices

4.2.1 Trusted QA Device

The Trusted QA Device forms an integral part of the system itself and resides within the trusted environment of the system. It enables the system to extend trust to external QA Device s. The Trusted QA Device is only trusted because the system itself is trusted.

4.2.2 External Untrusted QA Device

The External untrusted QA Device is a QA Device that resides external to the trusted environment of the system and is therefore untrusted. The purpose of the QA Chip Logical Interface is to allow the external untrusted QA Devices to become effectively trusted. This is accomplished when a Trusted QA Device shares a secret key with the external untrusted QA Device, or with a Translation QA Device (see below).

In a printing application external untrusted QA Devices would typically be instances of SBR4320 implementations located in a consumable or the printer.

4.2.3 Translation QA Device

A Translation QA Device is used to translate signatures between QA Devices and extend effective trust when secret keys are not directly shared between QA Devices.

The Translation QA Device must share a secret key with the Trusted QA Device that allows the Translation QA Device to effectively become trusted by the Trusted QA Device and hence trusted by the system. The Translation QA Device shares a different secret key with another external untrusted QA Device (which may in fact be a Translation QA Device etc). Although the Trusted QA Device doesn't share (know) the key of the external untrusted QA Device, signatures generated by that untrusted device can be translated by the Translation QA Device into signatures based on the key that the Trusted QA Device does know, and thus extend trust to the otherwise untrusted external QA Device.

In a SoPEC-based printing application, the Printer QA Device acts as a Translation QA Device since it shares a secret key with the SoPEC, and a different secret key with the ink carridges.

4.2.4 Consumable QA Device

A Consumable QA Device is an external untrusted QA Device located in a consumable. It typically contains details about the consumable, including how much of the consumable remains.

In a printing application the consumable QA Device is typically found in an ink cartridge and is referred to as an Ink QA Device, or simply Ink QA since ink is the most common consumable for printing applications. However, other consumables in printing applications include media and impression counts, so consumable QA Device is more generic.

4.2.5 Printer QA Device

A Printer QA Device is an external untrusted device located in the printer. It contains details about the operating parameters for the printer, and is often referred to as a Printer QA.

4.2.6 Value Upgrader QA Device

A Value Upgrader QA Device contains the necessary functions to allow a system to write an initial value (e.g. an ink amount) into another QA Device, typically a consumable QA Device. It also allows a system to refill/replenish a value in a consumable QA Device after use.

Whenever a value upgrader QA Device increases the amount of value in another QA Device, the value in the value upgrader QA Device is correspondingly decreased. This means the value upgrader QA Device cannot create value—it can only pass on whatever value it itself has been issued with. Thus a value upgrader QA Device can itself be replenished or topped up by another value upgrader QA Device.

An example of a value upgrader is an Ink Refill QA Device, which is used to fill/refill ink amount in an Ink QA Device.

4.2.7 Parameter Upgrader QA Device

A Parameter Upgrader QA Device contains the necessary functions to allow a system to write an initial parameter value (e.g. a print speed) into another QA Device, typically a printer QA Device. It also allows a system to change that parameter value at some later date.

A parameter upgrader QA Device is able to perform a fixed number of upgrades, and this number is effectively a consumable value. Thus the number of available upgrades decreases by 1 with each upgrade, and can be replenished by a value upgrader QA Device.

4.2.8 Key Programmer QA Device

Secret batch keys are inserted into QA Devices during instantiation (e.g. manufacture). These keys must be replaced by the final secret keys when the purpose of the QA Device is known. The Key Programmer QA Device implements all necessary functions for replacing keys in other QA Devices.

4.3 Signature

Digital signatures are used throughout the authentication protocols of the QA Chip Logical Interface.

A signature is produced by passing data plus a secret key through a keyed hash function. The signature proves that the data was signed by someone who knew the secret key.

The signature function used throughout the QA Chip Logical Interface is HMAC-SHA1 [1].

4.3.4 Authenticated Read

This is a read of data from a non-trusted QA Device that also includes a check of the signature (see Section 4.3.3). When the System determines that the signature is correct for the returned data (e.g. by asking a trusted QA Device to test the signature) then the System is able to trust that the data has not been tampered en route from the read, and was actually stored on the non-trusted QA Device.

4.3.5 Authenticated Write

An authenticated write is a write to the data storage area in a QA Device where the write request includes both the new data and a signature. The signature is based on a key that has write access permissions to the region of data in the QA Device, and proves to the receiving QA Device that the writer has the authority to perform the write. For example, a Value Upgrader Refilling Device is able to authorize a system to perform an authenticated write to upgrade a Consumable QA Device (e.g. to increase the amount of ink in an Ink QA Device).

The QA Device that receives the write request checks that the signature matches the data (so that it hasn't been tampered with en route) and also that the signature is based on the correct authorization key.

An authenticated write can be followed by an authenticated read to ensure (from the system's point of view) that the write was successful.

4.3.6 Non-Authenticated Write

A non-authenticated write is a write to the data storage area in a QA Device where the write request includes only the new data (and no signature). This kind of write is used when the system wants to update areas of the QA Device that have no access-protection.

The QA Device verifies that the destination of the write request has access permissions that permit anyone to write to it. If access is permitted, the QA Device simply performs the write as requested. A non-authenticated write can be followed by an authenticated read to ensure (from the system's point of view) that the write was successful.

4.3.7 Authorized Modification of Data

Authorized modification of data refers to modification of data via authenticated writes (see Section 4.3.5).

TABLE 2 List of data structures Group Represented description Name by Size Description QA Device Chip Identifier ChipId 48 bits Unique identifier for this QA Device. instance identifier Key and key Number of Keys NumKeys 8 Number of key slots available in this QA Device. related data Key K 160 bits K is the secret key used for calculating signatures. per key K^(n) is the key stored in the nth key slot. Key Identifier KeyId 31 bits Unique identifier for each key per key KeyId^(n) is the key identifier for the key stored in slot n. KeyLock KeyLock 1 bit Flag indicates whether the key is locked in the corresponding per key slot or not. KeyLock^(n) is the key lock flag for slot n. Operating and Number of NumVectors 4 Number of 512 bit memory vectors in this QA Device. state data Memory Vectors Memory Vector M 512 bits M is a 512 bit memory vector. per M¹ The 512-bit vector is divided into 16 × 32 bit words. M⁰ M⁰ stores application specific data that is protected by access permissions for key-based and non-key based writes. M¹ M¹ stores the attributes for M⁰, and is write-once-only. M²⁺ M²⁺ stores application specific data that is protected only by non key-based access permissions. Permissions P^(n) 16 bits Access permissions for each word of M¹⁺. n = number of M¹⁺ per P vectors Session data Random Number R 160 bits Current random number used to ensure time varying messages. Changes after each successful authentication or signature generation. 6 Instance/Device Identifier

Each QA Device requires an identifier that allows unique identification of that QA Device by external systems, ensures that messages are received by the correct QA Device, and ensures that the same device can be used across multiple transactions.

Strictly speaking, the identifier only needs to be unique within the context of a key, since QA Devices only accept messages that are appropriately signed. However it is more convenient to have the instance identifier completely unique, as is the case with this design.

The identifier functionality is provided by ChipId.

6.1 Chipid

ChipId is the unique 64-bit QA Device identifier. The ChipId is set when the QA Device is instantiated, and cannot be changed during the lifetime of the QA Device.

A 64-bit ChipId gives a maximum of 1844674 trillion unique QA Devices.

7 Key and Key Related Data

7.1 Numkeys, K, Keyid, And Keylock

Each QA Device contains a number of secret keys that are used for signature generation and verification. These keys serve two basic functions:

-   -   For reading, where they are used to verify that the read data         came from the particular QA Device and was not altered en route.     -   For writing, where they are used to ensure only authorised         modification of data.

Both of these functions are achieved by signature generation; a key is used to generate a signature for subsequent transmission from the device, and to generate a signature to compare against a received signature.

The number of secret keys in a QA Device is given by NumKeys. For this version of the QA Chip Logical Interface, NumKeys has a maximum value of 8.

Each key is referred to as K, and the subscripted form K_(n) refers to the nth key where n has the range 0 to NumKeys-1 (i.e. 0 to 7). For convenience we also refer to the nth key as being the key in the nth keyslot.

The length of each key is 160-bits. 160-bits was chosen because the output signature length from the signature generation function (HMAC-SHA1) is 160 bits, and a key longer than 160-bits does not add to the security of the function.

The security of the digital signatures relies upon keys being kept secret. To safeguard the security of each key, keys should be generated in a way that is not deterministic. Ideally each key should be programmed with a physically generated random number, gathered from a physically random phenomenon. Each key is initially programmed during QA Device instantiation.

Since all keys must be kept secret and must never leave the QA Device, each key has a corresponding 31-bit KeyId which can be read to determine the identity or label of the key without revealing the value of the key itself. Since the relationship between keys and KeyIds is 1:1, a system can read all the KeyIds from a QA Device and know which keys are stored in each of the keyslots.

Finally, each keyslot has a corresponding 1-bit KeyLock status indicating whether the key in that slot/position is allowed to be replaced (securely replaced, and only if the old key is known). Once a key has been locked into a slot, it cannot be unlocked i.e. it is the final key for that slot. A key can only be used to perform authenticated writes of data when it has been locked into its keyslot (i.e. its KeyLock status=1). Refer to Section 8.1.1.5 for further details.

Thus each of the NumKeys keyslots contains a 160-bit key, a 31-bit KeyID, and a 1-bit KeyLock.

7.2 Common And Variant Signature Generation

To create a digital signature, we pass the data to be signed together with a secret key through a key dependent one-way hash function. The key dependent one-way hash function used throughout the QA Chip Logical Interface is HMAC-SHA1[1].

Signatures are only of use if they can be validated i.e. QA Device A produces a signature for data and QA Device B can check if the signature was valid for that particular data. This implies that A and B must share some secret information so that they can generate equivalent signatures.

Common key signature generation is when QA Device A and QA Device B share the exact same key i.e. key K_(A)=key K_(B). Thus the signature for a message produced by A using K_(A) can be equivalently produced by B using K_(B). In other words SIG_(KA)(message)=SIG_(KB)(message) because key K_(A)=key K_(B).

Variant key signature generation is when QA Device B holds a base key, and QA Device A holds a variant of that key such that K_(A)=owf(K_(B),U_(A)) where owf is a one-way function based upon the base key (K_(B)) and a unique number in A (U_(A)). Thus A can produce SIG_(KA)(message), but for B to produce an equivalent signature it must produce K_(A) by reading U_(A) from A and using its base key K_(B). K_(A) is referred to as a variant key and K_(B) is referred to as the base/common key. Therefore, B can produce equivalent signatures from many QA Devices, each of which has its own unique variant of K_(B). Since ChipId is unique to a given QA Device, we use that as U_(A). A one-way function is required to create K_(A) from K_(B) or it would be possible to derive K_(B) if K_(A) were exposed.

Common key signature generation is used when A and B are equally available¹ to an attacker. For example, Printer QA Devices and Ink QA Devices are equally available to attackers (both are commonly available to an attacker), so shared keys between these two devices should be common keys. ¹The term “equally available” is relative. It typically means that the ease of availability of both are the effectively the same, regardless of price (e.g. both A and B are commercially available and effectively equally easy to come by).

Variant key signature generation is used when B is not readily available to an attacker, and A is readily available to an attacker. If an attacker is able to determine K_(A), they will not know K_(A) for any other QA Device of class A, and they will not be able to determine K_(B).

The QA Device producing or testing a signature needs to know if it must use the common or variant means of signature generation. Likewise, when a key is stored in a QA Device, the status of the key (whether it is a base or variant key) must be stored along with it for future reference. Both of these requirements are met using the KeyId as follows:

The 31-bit KeyId is broken into two parts:

-   -   A 30-bit unique identifier for the key. Bits 30-1 represents the         Id.     -   A 1-bit Variant Flag, which represents whether the key is a base         key or a variant key. Bit 0. represents the Variant Flag.

Table 247 describes the relationship of the Variant Flag with the key.

TABLE 247 Variant Flag representation value Key represented 0 Base key 1 Variant key 7.2.1 Equivalent Signature Generation Between QA Devices

Equivalent signature generation between 4 QA Devices A, B, C and D is shown in FIG. 363. Each device has a single key. KeyId.Id of all four keys are the same i.e KeyId_(A).Id=KeyId_(B).Id=KeyId_(C).Id=KeyId_(D).Id.

-   If KeyId_(A). VariantFlag=0 and KeyId_(B). VariantFlag=0, then a     signature produced by A, can be equivalently produced by B because     K_(A)=K_(B). -   If KeyId_(B). VariantFlag=0 and KeyId_(C). VariantFlag=1, then a     signature produced by C, is equivalently produced by B because     K_(C)=f (K_(B), ChipId_(C)). -   If KeyId_(C). VariantFlag=1 and KeyId_(D). VariantFlag=1, then a     signature produced by C, cannot be equivalently produced by D     because there is no common base key between the two devices. -   If KeyId_(D). VariantFlag=1 and KeyId_(A). VariantFlag=0, then a     signature produced by D, can be equivalently produced by A because     K_(D)=f (K_(A), ChipId_(D)).     8 Operating and State Data

The primary purpose of a QA Device is to securely hold application-specific data. For example if the QA Device is an Ink QA Device it may store ink characteristics and the amount of ink-remaining. If the QA Device is a Printer QA Device it may store the maximum speed and width of printing.

For secure manipulation of data:

-   -   Data must be clearly identified (includes typing of data).     -   Data must have clearly defined access criteria and permissions.

The QA Chip Logical Interface contains structures to permit these activities.

The QA Device contains a number of kinds of data with differing access requirements:

-   -   Data that can be decremented by anyone, but only increased in an         authorised fashion e.g. the amount of ink-remaining in an ink         cartridge.     -   Data that can only be decremented in an authorised fashion e.g.         the number of times a Parameter Upgrader QA Device has upgraded         another QA Device.     -   Data that is normally read-only, but can be written to (changed)         in an authorised fashion e.g. the operating parameters of a         printer.     -   Data that is always read-only and doesn't ever need to be         changed e.g. ink attributes or the serial number of an ink         cartridge or printer.     -   Data that is written by QACo/Silverbrook, and must not be         changed by the OEM or end user e.g. a licence number containing         the OEM's identification that must match the software in the         printer.     -   Data that is written by the OEM and must not be changed by the         end-user e.g. the machine number that filled the ink cartridge         with ink (for problem tracking).         8.1 M

M is the general term for all of the memory (or data) in a QA Device. M is further subscripted to refer to those different parts of M that have different access requirements as follows:

-   -   M₀ contains all of the data that is protected by access         permissions for key-based (authenticated) and non-key-based         (non-authenticated) writes.     -   M₁ contains the type information and access permissions for the         M₀ data, and has write-once permissions (each sub-part of M₁ can         only be written to once) to avoid the possibility of changing         the type or access permissions of something after it has been         defined.     -   M₂, M₃ etc., referred to as M₂₊, contains all the data that can         be updated by anyone until the permissions for those sub-parts         of M₂₊ have changed from read/write to read-only.

-   While all QA Devices must have at least M₀ and M₁, the exact number     of memory vectors (M_(n)s) available in a particular QA Device is     given by NumVectors. In this version of the QA Chip Logical     Interface there are exactly 4 memory vectors, so NumVectors=4.

-   Each M_(n) is 512 bits in length, and is further broken into 16×32     bit words. The ith word of M_(n) is referred to as M_(n)[i].     M_(n)[0] is the least significant word of M_(n), and M_(n)[15] is     the most significant word of M_(n).     8.1.1 M₀ and M₁

In the general case of data storage, it is up to the external accessor to interpret the bits in any way it wants. Data structures can be arbitrarily arranged as long as the various pieces of software and hardware that interpret those bits do so consistently. However if those bits have value, as in the case of a consumable, it is vital that the value cannot be increased without appropriate authorisation, or one type of value cannot be added to another incompatible kind e.g. dollars should never be added to yen.

Therefore M₀ is divided into a number of fields, where each field has a size, a position, a type and a set of permissions. M₀ contains all of the data that requires authenticated write access (one data element per field), and M₁ contains the field information i.e. the size, type and access permissions for the data stored in M₀.

Each 32-bit word of M₁ defines a field. Therefore there is a maximum of 16 defined fields. M₁[0] defines field 0, M₁[1] defines field 1 and so on. Each field is defined in terms of:

-   -   size and position, to permit external accessors determine where         a data item is     -   type, to permit external accessors determine what the data         represents     -   permissions, to ensure approriate access to the field by         external accessors.

The 32-bit value M₁[n] defines the conceptual field attributes for field n as follows: With regards to consistency of interpretation, the type, size and position information stored in the various words of M₁ allows a system to determine the contents of the corresponding fields (in M₀) held in the QA Device. For example, a 3-color ink cartridge may have an Ink QA Device that holds the amount of cyan ink in field 0, the amount of magenta ink in field 1, and the amount of yellow ink in field 2, while another single-color Ink QA Device may hold the amount of yellow ink in field 0, where the size of the fields in the two Ink QA Devices are different.

A field must be defined (in M₁) before it can be written to (in M₀). At QA Device instantiation, the whole of M₀ is 0 and no fields are defined (all of M₁ is 0). The first field (field 0) can only be created by writing an appropriate value to M₁[0]. Once field 0 has been defined, the words of M₀ corresponding to field 0 can be written to (via the appropriate permissions within the field definition M₁[0]).

Once a field has been defined (i.e. M₁[n] has been written to), the size, type and permissions for that field cannot be changed i.e. M₁ is write-once. Otherwise, for example, a field could be defined to be lira and given an initial value, then the type changed to dollars.

The size of a field is measured in terms of the number of consecutive 32-bit words it occupies.

Since there are only 16×32-bit words in M₀, there can only be 16 fields when all 16 fields are defined to be 1 word sized each. Likewise, the maximum size of a field is 512 bits when only a single field is defined, and it is possible to define two fields of 256-bits each.

Once field 0 has been created, field 1 can be created, and so on. When enough fields have been created to allocate all of M₀, the remaining words in M₁ are available for write-once general data storage purposes.

It must be emphasised that when a field is created the permissions for that field are final and cannot be changed. This also means that any keys referred to by the field permissions must be already locked into their keyslots. Otherwise someone could set up a field's permissions that the key in a particular keyslot has write access to that field without any guarantee that the desired key will be ever stored in that slot (thus allowing potential mis-use of the field's value).

8.1.1.1 Field Size and Position

A field's size and position are defined by means of 4 bits (referred to as EndPos) that point to the least significant word of the field, with an implied position of the field's most significant word. The implied position of field 0's most significant word is M₀[15]. The positions and sizes of all fields can therefore be calculated by starting from field 0 and working upwards until all the words of M₀ have been accounted for.

The default value of M₁[0] is 0, which means field0.endPos=0. Since field0.startPos=15, field 0 is the only field and is 16 words long.

8.1.1.1.1 EXAMPLE

Suppose for example, we want to allocate 4 fields as follows:

-   -   field 0: 128 bits (4×32-bit words)     -   field 1: 32 bits (1×32-bit word)     -   field 2: 160 bits (5×32-bit words)     -   field 3: 192 bits (6×32-bit words)

Field 0's position and size is defined by M₁[0], and has an assumed start position of 15, which means the most significant word of field 0 must be in M₀[15]. Field 0 therefore occupies M₀[12] through to M₀[15], and has an endPos value of 12.

Field 1's position and size is defined by M₁[1], and has an assumed start position of 11 (i.e. M₁[0].endPos-1). Since it has a length of 1 word, field 1 therefore occupies only M₀[11] and its end position is the same as its start position i.e. its endPos value is 11.

Likewise field 2's position and size is defined by M₁[2], and has an assumed start position of 10 (i.e. M₁[1].endPos-1). Since it has a length of 5 words, field 2 therefore occupies M₀[6] through to M₀[10] and and has an endPos value of 6.

Finally, field 3's position and size is defined by M₁[3], and has an assumed start position of 5 (i.e. M₁[2].endPos-1). Since it has a length of 6 words, field 3 therefore occupies M₀[5] through to M₀[0] and and has an endPos value of 0.

Since all 16 words of M₀ are now accounted for in the 4 fields, the remaining words of M₁ (i.e. M₁[4] though to M₁[15]) are ignored, and can be used for any write-once (and thence read-only) data.

FIG. 365 shows the same example in diagramatic format.

8.1.1.1.2 Determining the Number of Fields

The following pseudocode illustrates a means of determining the number of fields:

fieldNum FindNumFields(M1) startPos

15 fieldNum

0 While (fieldNum < 16) endPos

M1[fieldNum].endPos If (endPos > startPos) # error in this field... so must be an attack attackDetected( ) # most likely clears all keys and data EndIf fieldNum++ If (endPos = 0) return fieldNum # is already incremented Else startPos

endPos − 1 # endpos must be > 0 EndIf EndWhile # error if get here since 16 fields are consumed in 16 words at most attackDetected( ) # most likely clears all keys and data 8.1.1.1.3 Determining the Sizes of All Fields

The following pseudocode illustrates a means of determing the sizes of all valid fields:

FindFieldSizes(M1, fieldSize[ ]) numFields

FindNumFields (M1) # assumes that FindNumFields does all checking ntartPos

15 fieldNum

0 While (fieldNum < numFields) EndPos

M1[fieldNum].endPos fieldSize[fieldNum] = startPos − endPos + 1 startPos

endPos − 1  # endpos must be > 0 fieldNum++ EndWhile While (fieldNum < 16) fieldSize[fieldNum]

0 fieldNum++ EndWhile 8.1.1.2 Field Type

The system must be able to identify the type of data stored in a field so that it can perform operations using the correct data. For example, a printer system must be able identify which of a consumable's fields are ink fields (and which field is which ink) so that the ink usage can be correctly applied during printing.

A field's type is defined by 15 bits. Table 332 in Appendix A lists the field types that are specifically required by the QA Chip Logical Interface and therefore apply across all applications.

The default value of M₁[0] is 0, which means field0.type=0 (i.e. non-initialised).

Strictly speaking, the type need only be interpreted by all who can securely read and write to that field i.e. within the context of one or more keys. However it is convenient if possible to keep all types unique for simplistic identification of data across all applications.

In the general case, an external system communicating with a QA Device can identify the data stored in M0 in the following way:

-   -   Read the KeyId of the key that has permission to write to the         field. This will a give broad identification of the data type,         which may be sufficient for certain applications.     -   Read the type attribute for the field to narrow down the         identity within the broader context of the KeyId.

For example, the printer system can read the KeyId to deduce that the data stored in a field can be written to via the HP_Network_InkRefill key, which means that any data is of the general ink category known to HP Network printers. By further reading the type attribute for the field the system can determine that the ink is Black ink.

8.1.1.3 Field Permissions

All fields can be ready by everyone. However writes to fields are governed by 13-bits of permissions that are present in each field's attribute definition. The permissions describe who can do what to a specific field.

Writes to fields can either be authenticated (i.e. the data to be written is signed by a key and this signature must be checked by the receiving device before write access is given) or non-authenticated (i.e. the data is not signed by a key). Therefore we define a single bit (AuthRW) that specifies whether authenticated writes are permitted, and a single bit (NonAuthRW) specifying whether non-authenticated writes are permitted. Since it is pointless to permit both authenticated and non-authenticated writes to write any value (the authentciated writes are pointless), we further define the case when both bits are set to be interpreted as authenticated writes are permitted, but non-authenticated writes only succeed when the new value is less than the previous value i.e. the permission is decrement-only. The interpretation of these two bits is shown in Table 249.

TABLE 249 Interpretation of AuthRW and NonAuthRW NonAuthRW AuthRW Interpretation 0 0 Read-only access (no-one can write to this field). This is the initial state for each field. At instantiation all of M₁ is 0 which means AuthRW and NonAuthRW are 0 for each field, and hence none of M₀ can be written to until a field is defined. 0 1 Authenticated write access is permitted Non-authenticated write acecss is not permitted 1 0 Authenticated write access is not permitted Non-authenticated write access is permitted (i.e. anyone can write to this field) 1 1 Authenticated write access is permitted Non-authenticated write access is decrement-only.

If authenticated write access is permitted, there are 11 additional bits (bringing the total number of permission bits to 13) to more fully describe the kind of write access for each key. We only permit a single key to have the ability to write any value to the field, and the remaining keys are defined as being either not permitted to write, or as having decrement-only write access. A 3-bit KeyNum represents the slot number of the key that has the ability to write any value to the field (as long as the key is locked into its key slot), and an 8-bit KeyPerms defines the write permissions for the (maximum of) 8 keys as follows:

-   -   KeyPerms[n]=0: The key in slot n (i.e. K_(n)) has no write         access to this field (except when n=KeyNum). Setting KeyPerms to         0 prohibits a key from transferring value (when an amount is         deducted from field in one QA Device and transferred to another         field in a different QA Device)     -   KeyPerms[n]=1: The key in slot n (i.e. K_(n)) is permitted to         perform decrement-only writes to this field (as long as K_(n) is         locked in its key slot). Setting KeyPerms to 1 allows a key to         transfer value (when an amount is deducted from field in one QA         Device and transferred to another field in a different QA         Device).

The 13-bits of permissions (within bits 4-16 of M₁[n]) are allocated as follows:

8.1.1.3.1 Example 1

-   -   FIG. 367 shows an example of permission bits for a field.     -   In this example we can see:     -   NonAuthRW=0 and AuthRW=1, which means that only authenticated         writes are allowed i.e. writes to the field without an         appropriate signature are not permitted.     -   KeyNum=3, so the only key permitted to write any value to the         field is key 3 (i.e. K₃).     -   KeyPerms[3]=0, which means that although key 3 is permitted to         write to this field, key 3 can't be used to transfer value from         this field to other QA Devices.     -   KeyPerms[0,4,5,6,7]=0, which means that these respective keys         cannot write to this field.     -   KeyPerms[1,2]=1, which means that keys 1 and 2 have         decrement-only access to this field i.e. they are permitted to         write a new value to the field only when the new value is less         than the current value.         8.1.1.3.2 Example 2     -   FIG. 368 shows a second example of permission bits for a field.     -   In this example we can see:     -   NonAuthRW and AuthRW=1, which means that authenticated writes         are allowed and writes to the field without a signature are only         permitted when the new value is less than the current value         (i.e. non-authenticated writes have decrement-only permission).     -   KeyNum=3, so the only key permitted to write any value to the         field is key 3 (i.e. K₃).     -   KeyPerms[3]=1, which means that key 3 is permitted to write to         this field, and can be used to transfer value from this field to         other QA Devices.     -   KeyPerms[0,4,5,6,7]=0, which means that these respective keys         cannot write to this field.     -   KeyPerms[1,2]=1, which means that keys 1 and 2 have         decrement-only access to this field i.e. they are permitted to         write a new value to the field only when the new value is less         than the current value.         8.1.1.4 Summary of Field Attributes     -   FIG. 369 shows the breakdown of bits within the 32-bit field         attribute value M₁[n].     -   Table 250 summarises each attribute.

TABLE 250 Attributes for a field Sub- attribute Size Attribute name in bits Interpretation Type Type 15 Gives additional identification of the data stored in the field within the context of the accessors of that field. Permissions KeyNum 3 The slot number of the key that has authenticated write access to the field. NonAuthRW 1 0 = non-authenticated writes are not permitted to this field. 1 = non-authenticated writes are permitted to this field (see Table 249). AuthRW 1 0 = authenticated writes are not permitted to this field. 1 = authenticated writes are permitted to this field. KeyPerms 8 Bitmap representing the write permissions for each of the keys when AuthRW = 1. For each bit: 0 = no write access for this key (except for key KeyNum) 1 = decrement-only access is permitted for this key. Size and EndPos 4 The word number in M₀ that holds Position the lsw of the field. The msw is held in M1[fieldNum − 1], where msw of field 0 is 15. 8.1.1.5 Permissions of M₁

M₁ holds the field attributes for data stored in M₀, and each word of M₁ can be written to once only.

It is important that a system can determine which words are available for writing. While this can be determined by reading M₁ and determining which of the words is non-zero, a 16-bit permissions value P₁ is available, with each bit indicating whether or not a given word in M₁ has been written to. Bit n of P₁ represents the permissions for M₁[n] as follows:

TABLE 251 Interpretation of P₁[n] i.e. bit n of M₁'s permission Description 0 writes to M₁[n] are not permitted i.e. this word is now read-only 1 writes to M₁[n] are permitted

Since M₁ is write-once, whenever a word is written to in M₁, the corresponding bit of P₁ is also cleared, i.e. writing to M₁[n] clears P₁[n].

Writes to M₁[n] only succeed when all of M₁[0...n−1] have already written to (i.e. previous fields are defined) i.e.

-   -   M₁[0..n−1] must have already been written to (i.e. P₁[0..n−1]         are 0)     -   P₁[n]=1 (i.e. it has not yet been written to)

In addition, if M₁[n−1].endPos≠0, the new M₁[n] word will define the attributes of field n, so must be further checked as follows:

-   -   The new M₁[n].endPos must be valid (i.e. must be less than         M₁[n−1].endPos)     -   If the new M₁[n].authRW is set, K_(keyNum) must be locked, and         all keys referred to by the new M₁[n].keyPerms must also be         locked.

However if M₁[n−1].endPos=0, then all of M₀ has been defined in terms of fields. Since enough fields have been created to allocate all of M₀, any remaining words in M₁ are available for write-once general data storage purposes, and are not checked any further.

8.1.2 M2+

M₂, M₃ etc., referred to as M₂+, contains all the data that can be updated by anyone (i.e. no authenticated write is required) until the permissions for those sub-parts of M₂₊ have changed from read/write to read-only.

The same permissions representation as used for M₁ is also used for M₂₊. Consequently P_(n) is a 16-bit value that contains the permissions for M_(n) (where n>0). The permissions for word w of M_(n) is given by a single bit P_(n)[w]. However, unlike writes to M₁, writes to M₂₊ do not automatically clear bits in P. Only when the bits in P₂₊ are explictly cleared (by anyone) do those corresponding words become read-only and final.

9 Session Data

Data that is valid only for the duration of a particular communication session is referred to as session data. Session data ensures that every signature contains different data (sometimes referred to as a nonce) and this prevents replay attacks.

9.1 R

R is a 160-bit random number seed that is set up (when the QA Device is instantiated) and from that point on it is internally managed and updated by the QA Device. R is used to ensure that each signed item contains time varying information (not chosen by an attacker), and each QA Device's R is unrelated from one QA Device to the next.

This R is used in the generation and testing of signatures.

An attacker must not be able to deduce the values of R in present and future devices. Therefore, R should be programmed with a cryptographically strong random number, gathered from a physically random phenomenon (must not be deterministic).

9.2 Advancing R

The session component of the message must only last for a single session (challenge and response).

The rules for updating R are as follows:

-   -   Reads of R do not advance R.     -   Everytime a signature is produced with R, R is advanced to a new         random number.     -   Everytime a signature including R is tested and is found to be         correct, R is advanced to a new random number.         9.3 R_(L) and R_(E)

Each signature contains 2 pieces of session data i.e. 2 Rs:

-   -   One R comes from the QA Device issuing the challenge i.e. the         challenger. This is so the challenger can ensure that the         challenged QA Device isn't simply replaying an old signature         i.e. the challenger is protecting itself against the challenged.     -   One R comes from the device responding to the challenge i.e. the         challenged. This is so the challenged never signs anything that         is given to it without inserting some time varying change i.e.         protects the challenged from the challenger in case the         challenger is actually an attacker performing a chosen text         attack

Since there are two Rs, we need to distinguish between them. We do so by defining each R as external (R_(E)) or local (R_(L)) depending on its use in a given function. For example, the challenger sends out its local R, referred to as R_(L). The device being challenged receives the challenger's R as an external R, i.e R_(E). It then generates a signature using its R_(L) and the challenger's R_(E). The resultant signature and R_(L) are sent to the challenger as the response. The challenger receives the signature and R_(E) (signature and R_(L) produced by the device being challenged), produces its own signature using R_(L) (sent to the device being challenged earlier) and R_(E) received, and compares that signature to the signature received as response.

Signature Functions

10 Objects

10.1 KeyRef

10.1.1 Object Description

Instead of passing keys directly into a function, a KeyRef (i.e. key reference) object is passed instead. A KeyRef object encapsulates the process by which a key is formed for common and variant forms of signature generation (based on the setting of the variables within the object). A KeyRef defines which key to use, whether it is a common or variant form of that key, and, if it is a variant form, the ChipId to use to create the variant. For more information about common and variant forms of keys, see Section 7.2.

Users pass KeyRef objects in as input parameters to public functions of the QA Chip Logical Interface, and these KeyRefs are subsequently passed to the signature function (called within the interface function). Note, however, that the method functions for KeyRef objects are not available outside the QA Chip Logical Interface.

10.1.2 Object Variables

Table 252 describes each of the variables within a KeyRef object.

TABLE 252 Description of object variables for KeyRef object Parameter Description keyNum Slot number of the key to use as the basis for key formation useChipId 0 = the key to be formed is a common key (i.e. is the same as K_(keyNum)) 1 = the key to be formed is a variant key based on K_(keyNum) ChipId When useChipId = 1, this is the ChipId to be used to form the variant key (this will be the ChipId of the QA Device which stores the variant of K_(keyNum)) When useChipId = 0, chipId is not used 10.1.3 Object Methods 10.1.3.1 getKey

-   -   public key getKey(void)         10.1.3.1.1 Method Description

This method is a public method (public in object oriented terms, not public to users of the QA Chip Logical Interface) and is called by the GenerateSignature function to return the key for use in signature generation.

If useChipId is true, the formKeyVariant method is called to form the key using chipid and then return the variant key. If useChipId is false, the key stored in slot keyNum is returned.

10.1.3.1.2 Method Sequence

The getKey method is illustrated by the following pseudocode:

-   -   If (useChipId=0)         -   key←K_(keyNum)     -   Else         -   key←formKeyVariant()     -   EndIf     -   Return key         10.1.3.2 formKeyVariant     -   private key formKeyVariant (voID)         10.1.3.2.1 Method Description

This method produces the variant form of a key, based on the K_(keyNum) and chipId. As described in Section 7.2, the variant form of key K_(keyNum) is generated by owf (K_(keyNum), chipID) where owf is a one-way function.

In addition, the time taken by owf must not depend on the value of the key i.e. the timing should be effectively constant. This prevents timing attacks on the key.

At present, owf is SHA1, although this still needs to be verified. Thus the variant key is defined to be SHA1(K_(keyNum)|chipID).

10.1.3.2.2 Method Sequence

The formKeyVariant method is illustrated by the following pseudocode:

-   -   key←SHA1(K_(keyNum)|chipId)# Calculation must take constant time     -   Return key         11 Functions

Digital signatures form the basis of all authentication protocols within the QA Chip Logical Interface. The signature functions are not directly available to users of the QA Chip Logical Interface, since a golden rule of digital signatures is never to sign anything exactly as it has been given to you.

Instead, these signature functions are internally available to the functions that comprise the public interface, and are used by those functions for the formation of keys and the generation of signatures.

11.1 GENERATESIGNATURE

-   -   Input: KeyRef, Data, Random1, Random2     -   Output SIG     -   Changes: None     -   Availability: All devices         11.1.1 Function Description

This function uses KeyRef to obtain the actual key required for signature generation, appends Random1 and Random2 to Data, and performs HMAC_SHA1[key, Data] to output a signature. HMAC_SHA1 is described in [1]. In addition, this operation must take constant time irrespective of the value of the key (see Section 10.1.3.2 for more details).

11.1.2 Input Parameter Description

Table 253 describes each of the input parameters:

TABLE 253 Description of input parameters for GenerateSignature Parameter Description KeyRef This is an instance of the KeyRef object for use by the GenerateSignature function. For common key signature generation: KeyRef.keyNum = Slot number of the key to be used to produce the signature. KeyRef.useChipId = 0 For variant key signature generation: KeyRef.keyNum = Slot number of the key to be used for generating the variant key, where the variant key is to be used to produce the signature KeyRef.useChipId = 1 KeyRef.chipId = ChipId of the QA Device which stores the variant of K_(KeyRef.keyNum), and uses the variant key for signature generation Data Preformatted data to be signed. Random1 and Random2 are appended to Data before the signature is generated to ensure that the signature is session based (applicable only to a single session). Random1 This is the session component from the QA Device that is responding to the challenge. Random2 This is the session component from the QA Device that issued the challenge. 11.1.3 Output Parameter Description

Table 254 describes each of the output parameters.

TABLE 254 Description of output parameters for GenerateSignature Parameter Description SIG SIG = SIG_(key)(Data | Random1 | Random2) where key = KeyRef.getKey( ) 11.1.4 Function Sequence

The GenerateSignature function is illustrated by the following pseudocode:

-   -   key←KeyRef.getKey()     -   dataToBeSigned←Data|Random1|Random2     -   SIG←HMAC_SHA1(key, dataToBeSigned)# Calculation must take     -   constant time     -   Output SIG     -   Return         Basic Functions         12 Definitions

This section defines return codes and constants referred to by functions and pseudocode.

12.1 ResultFlag

The ResultFlag is a byte that indicates the return status from a function. Callers can use the value of ResultFlag to determine whether a call to a function succeeded or failed, and if the call failed, the specific error condition.

Table 255 describes the ResultFlag values and the mnemonics used in the pseudocode.

TABLE 255 ResultFlag value description Mnemonic Description Possible causes Pass Function Function successfully completed completed requested task. sucessfully Fail General An error occurred during Failure function processing. BadSig Signature Input signature didn't match mismatch the generated signature. InvalidKey KeyRef Input KeyRef.keyNum > 3. incorrect InvalidVector VectNum Input M_(VectNum) > 3. incorrect InvalidPermission Permission Trying to perform a Write or not adqeuate WriteAuth with incorrect to per form permissions. operation. KeyAlreadyLocked Key already Key cannot be changed because locked. it has already been locked. 12.2 Constants

Table 256 describes the constants referred to by functions and pseudocode.

TABLE 256 Constants Definition Value MaxKey NumKeys −1 (typically 7) MaxM NumVectors −1 (typically 3) MaxWordInM 16 − 1 = 15 13 GetInfo

-   -   Input: None     -   Output: ResuitFlag, SoftwareReleaseIdMajor,         SoftwareReleaseIdMinor,         -   NumVectors, NumKeys,ChipId         -   DepthOfRollBackCache (for an upgrade device only)     -   Changes: None     -   Availability: All devices         13.1 Function Description

Users of QA Devices must call the GetInfo function on each QA Device before calling any other functions on that device.

The GetInfo function tells the caller what kind of QA Device this is, what functions are available and what properties this QA Device has. The caller can use this information to correctly call functions with appropriately formatted parameters.

The first value returned, SoftwareReleaseIdMajor, effectively identifies what kind of QA Device this is, and therefore what functions are available to callers. SoftwareReleaseIdMinor tells the caller which version of the specific type of QA Device this is. The mapping between the SoftwareReleaseIdMajor and type of device and their different functions is described in Table 258 Every QA Device also returns NumVectors, NumKeys and ChipId which are required to set input parameter values for commands to the device.

Additional information may be returned depending on the type of QA Device. The VarDataLen and VarData fields of the output hold this additional information.

13.2 Output Parameters

Table 257 describes each of the output parameters.

TABLE 257 Description of output parameters for GetInfo function # Parameter bytes Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1. SoftwareReleaseIdMajor 1 This defines the function set that is available on this QA Device. SoftwareReleaseIdMinor 1 This defines minor software releases within a major release, and are incremental changes to the software mainly to deal with bug fixes. NumVectors 1 Total number of memory vectors in this QA Device. NumKeys 1 Total number of keys in this QA Device. ChipId 6 This QA Device's ChipId VarDataLen 1 Length of bytes to follow. VarData (VarDataLen This is additional bytes) application specific data, and will be of length VarDataLen (i.e. may be 0).

Table 258 shows the mapping between the SoftwareReleaseIdMajor, the type of QA Device and the available device functions.

TABLE 258 Mapping between SoftwareReleaseIdMajor and available device functions Software ReleaseId Major Device description Functions available 1 Ink or Printer QA GetInfo Device Random Read Test Translate WriteM1+ WriteFields WriteFieldsAuth SetPerm ReplaceKey 2 Value Upgrader QA All functions in Device (e.g. Ink the Ink or Printer Refill QA Device) Device, plus: StartXfer XferAmount StartRollBack RollBackAmount 3 Parameter Upgrader All functions in QA Device the Ink or Printer device, plus: StartXfer XferField StartRollBack RollBackField 4 Key Replacement All functions in device the Ink or Printer Device, plus: GetProgramKey ReplaceKey - is different from the Ink or Printer device 5 Trusted device All functions in the Ink or Printer Device, plus: SignM

Table 259 shows the VarData components for Value Upgrader and Parameter Upgrader QA Devices.

TABLE 259 VarData for Value and Parameter Upgrader QA Devices VarData Length in Components bytes Description DepthOfRollBackCache 1 The number of datasets that can be accommodated in the Xfer Entry cache of the device. 13.3 Function Sequence

The GetInfo command is illustrated by the following pseudocode:

-   -   Output SoftwareReleaseIdMajor     -   Output SoftwareReleaseIdMinor     -   Output NumVectors     -   Output NumKeys     -   Output ChipId     -   VarDataLen←1# In case of an upgrade device     -   Output DepthOfRollBackCache     -   Return         14 Random     -   Input: None     -   Output: R_(L)     -   Changes: None     -   Availability: All devices

The Random command is used by the caller to obtain a session component (challenge) for use in subsequent signature generation.

If a caller calls the Random function multiple times, the same output will be returned each time. R_(L) (i.e. this QA Device's R) will only advance to the next random number in the sequence after a successful test of a signature or after producing a new signature. The same R_(L) can never be used to produce two signatures from the same QA Device.

The Random command is illustrated by the following pseudocode:

-   -   Output: R_(L)     -   Return         15 Read     -   Input KeyRef, SigOnly, MSelect, KeyIdSelect, WordSelect, R_(E)     -   Output: ResultFlag, SelectedWordsOfSelectedMs, SelectedKeyIds,         R_(L), SIG_(out)     -   Changes: R_(L)     -   Availability: All devices         15.1 Function Description

The Read command is used to read data and keyIds from a QA Device. The caller can specify which words from M and which KeyIds are read.

The Read command can return both data and signature, or just the signature of the requested data.

Since the return of data is based on the caller's input request, it prevents unnecessary information from being sent back to the caller. Callers typically request only the signature in order to confirm that locally cached values match the values on the QA Device.

The data read from an untrusted QA Device (A) using a Read command is validated by a trusted QA Device (B) using the Test command. The R_(L) and SIG_(out) produced as output from the Read command are input (along with correctly formatted data) to the Test command on a trusted QA Device for validation of the signature and hence the data. SIG_(out) can also optionally be passed through the Translate command on a number of QA Devices between Read and Test if the QA Devices A and B do not share keys.

15.2 Input Parameters

Table 260 describes each of the input parameters:

TABLE 260 Description of input parameters for Read Parameter Description KeyRef For common key signature generation: KeyRef.keyNum = Slot number of the key to be used for producing the output signature. KeyRef.useChipId = 0 No variant key signature generation required SigOnly Flag indicating return of signature and data. 0- indicates both the signature and data are to be returned. 1- indicates only the signature is to be returned. Mselect Selection of memory vectors to be read - each bit corresponding to a given memory vector (a maximum of NumVector bits) 0- indicates the memory vector must not be read. 1- indicates memory vector must be read. KeyIdSelect Selection of KeyIds to be read - each bit corresponds to a given KeyId (a maximum of NumKey bits). 0- indicates KeyId must not be read. 1- indicates KeyId must be read. WordSelect Selection of words read from a desired M as requested in MSelect. Each WordSelect is 16 bits corresponding to each bit in MSelect. Each bit in the WordSelect indicates whether or not to read the corresponding word for the particular M. 0- indicates word must not be read. 1- indicates word must be read. R_(E) External random value required for output signature generation (i.e the challenge). R_(E) is obtained by calling the Random function on the device which will receive the SIG_(out) from the Read function. 15.3 Output Parameters

Table 261 describes each of the output parameters.

Parameter Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1. SelectedWordsOfSelectedMs Selected words from selected memory vectors as requested by MSelect and WordSelect. SelectedKeyIds Selected KeyIds as requested by KeyIdSelect. R_(L) Local random value added to the output signature (i.e SIG_(out)). Refer to FIG. 370. SIG_(out) SIG_(out) = SIG_(KeyRef)(data | R_(L) | R_(E)) as shown in FIG. 8. Refer to Section 10.1.3.1 for details 15.3.1 SIG_(out)

-   -   FIG. 370 shows the formatting of data for output signature         generation.     -   Table 262 gives the parametersincluded in SIG_(out)

Value set Length Value set from Parameter in bits internally Input RWSense 3 read constant = 000 Refer to Section 15.3.1.1 MSelect 4 ● KeyIdSelect 8 ● ChipId 48 This QA Device's ChipId WordSelect 16 ● per M SelectedWordsOfSelectedMs 32 The appropriate ● per word words from the various Ms as selected by the caller R_(L) 160 This QA Device's current R R_(E) 160 ● 15.3.1.1 RWSense

An RWSense value is present in the signed data to distinguish whether a signature was produced from a Read or produced for a WriteAuth.

The RWSense is set to a read constant (000) for producing a signature from a read function. The RWSense is set to a write constant (001) for producing a signature for a write function.

The RWSense prevents signatures produced by Read to be subsequently sent into a WriteAuth function. Only signatures produced with RWSense set to write (001), are accepted by a write function.

15.4 Function Sequence

The Read command is illustrated by the following pseudocode:

Accept input parameters- KeyRef, SigOnly, MSelect, KeyIdSelect # Accept input parameter WordSelect based on MSelect For i

0 to MaxM If(MSelect[i] = 1) Accept next WordSelect WordSelectTemp[i]

WordSelect EndIf EndFor Accept R_(E) Check range of KeyRef.keyNum If invalid ResultFlag

InvalidKey Output ResultFlag Return EndIf #Build SelectedWordsOfSelectedMs k

0 # k stores the word count for SelectedWordsOfSelectedMs SelectedWordsOfSelectedMs[k]

0 For i

0 to 3 If(MSelect[i] = 1) For j

0 to MaxWordInM If(WordSelectTemp[i][j] = 1)  SelectedWordsOfSelectedMs[k]

(M_(i)[j])  k++ EndIf EndFor EndIf EndFor #Build SelectedKeyIds l

0 # l stores the word count for SelectedKeyIds SelectedKeyIds[l]

0 For i

0 to MaxKey If(KeyIdSelect[i] = 1) SelectedKeyIds[l]

KeyId[i] l++ EndIf EndFor #Generate message for passing into the GenerateSignature function data

(RWSense|MSelect|KeyIdSelect|ChipId|WordSelect |SelectedWordsOfSelectedMs|SelectedKeylds) # Refer to Figure 370. #Generate Signature function SIG_(L)

GenerateSignature(KeyRef,data,R_(L),R_(E)) # See Section 11.1 Update R_(L) to R_(L2) ResultFlag

Pass Output ResultFlag If(SigOnly = 0) Output SelectedWordsOfSelectedMs, SelectedKeyIds EndIf Output R_(L), SIG_(L) Return 16 Test

-   -   Input: KeyRef, DataLength, Data, R_(E), SIG_(E)     -   Output: ResultFlag     -   Changes: R_(L)     -   Availability: All devices except ink device         16.1 Function Description

The Test command is used to validate data that has been read from an untrusted QA Device according to a digital signature SIG_(E). The data will typically be memory vector and KeyId data. SIG_(E) (and its related R_(E)) is the most recent signature—this will be the signature produced by Read if Translate was not used, or will be the output from the most recent Translate if Translate was used. The Test function produces a local signature (SIG_(L)=SIG_(key)(Data|R_(E)|R_(L)) and compares it to the input signature (SIG_(E)). If the two signatures match the function returns ‘Pass’, and the caller knows that the data read can be trusted.

The key used to produce SIG_(L) depends on whether SIG_(E) was produced by a QA Device sharing a common key or a variant key. The KeyRef object passed into the interface must be set appropriately to reflect this.

The Test function accepts preformatted data (as DataLength number of words), and appends the external R_(E) and local R_(L) to the preformatted data to generate the signature as shown in FIG. 371.

16.2 Input Parameters

Table 263 describes each of the input parameters.

TABLE 263 Description of input parameters for Test Parameter Description KeyRef For testing common key signature: KeyRef.keyNum = Slot number of the key to be used for testing the signature. SIG_(E) produced using K_(KeyRef.keyNum) by the external device. KeyRef.useChipId = 0 For testing variant key signature: KeyRef.keyNum = Slot number of the key to be used for generating the variant key. SIG_(E) produced using a variant of K_(KeyRef.keyNum) by the external device. KeyRef.useChipId = 1 KeyRef.chipId = ChipId of the device which generated SIG_(E) using a variant of K_(KeyRef.keyNum). DataLength Length of preformatted data in words. Must be non zero. Data Preformatted data to be used for producing the signature. R_(E) External random value required for verifying the input signature. This will be the R from the input signature generator (i.e the device generating SIG_(E)). SIG_(E) External signature required for authenticating input data as shown in FIG. 371. The external signature is generated either by a Read function or a Translate function. A correct SIG_(E) = SIG_(KeyRef)(Data | R_(E) | R_(L)). 16.2.1 Input Signature Verification Data Format

FIG. 371 shows the formatting of data for input signature verification.

The data in FIG. 371 (i.e. not R_(E) or R_(L)) is typically output from a Read function (formatted as per FIG. 370). The data may also be generated in the same format by the system from its cache as will be the case when it performs a Read using SigOnly=1.

16.3 Output Parameters

Table 264 describes each of the output parameters.

TABLE 264 Description of output parameters for Test Parameter Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1. 16.4 Function Sequence

The Test command is illustrated by the following pseudocode:

Accept input parameters-KeyRef, DataLength # Accept input parameter- Data based on DataLength For i

0 to (DataLength − 1) Accept next word of Data EndFor Accept input parameters - R_(E), SIG_(E) Check range of KeyRef.keyNum If invalid ResultFlag

InvalidKey Output ResultFlag Return EndIf #Generate signature SIG_(L)

GenerateSignature(KeyRef,Data,R_(E),R_(L)) # Refer to Figure 371. #Check signature If(SIG_(L) = SIG_(E)) Update R_(L) to R_(L2) ResultFlag

Pass Else ResultFlag

BadSig EndIf Output ResultFlag Return 17 Translate

-   Input InputKeyRef, DataLength, Data, R_(E), SIG_(E), OutputKeyRef,     R_(E2) -   Output: ResultFlag, R_(L2), SIG_(out) -   Changes: R_(L) -   Availability: Printer device, and possibly on other devices     17.1 Function Description

It is possible for a system to call the Read function on QA Device A to obtain data and signature, and then call the Test function on QA Device B to validate the data and signature. In the same way it is possible for a system to call the SignM function on a trusted QA Device B and then call the WriteAuth function on QA Device B to actually store data on B. Both of these actions are only possible when QA Devices A and B share secret key information.

If however, A and B do not share secret keys, we can create a validation chain (and hence extension of trust) by means of translation of signatures. A given QA Device can only translate signatures if it knows the key of the previous stage in the chain as well as the key of the next stage in the chain. The Translate function provides this functionality.

The Translate function translates a signature from one based on one key to one based another key. The Translate function first performs a test of the input signature using the InputKeyRef, and if the test succeeds produces an output signature using the OutputKeyRef. The Translate function can therefore in some ways be considered to be a combination of the Test and Read function, except that the data is input into the QA Device instead of being read from it.

The InputKeyRef object passed into Translate must be set appropriately to reflect whether SIG_(E) was produced by a QA Device sharing a common key or a variant key.

The key used to produce output signature SIG_(out) depends on whether the translating device shares a common key or a variant key with the QA Device receiving the signature. The OutputKeyRef object passed into Translate must be set appropriately to reflect this.

Since the Translate function does not interpret or generate the data in any way, only preformatted data can be passed in. The Translate function does however append the external R_(E) and local R_(L) to the preformatted data for verifying the input signature, then advances R_(L) to R_(L2), and appends R_(L2) and R_(E2) to the preformatted data to produce the output signature. This is done to protect the keys and prevent replay attacks.

The Translate Functions Translates:

-   -   signatures for subsequent use in Test, typically originating         from Read     -   signatures for subsequent use in WriteAuth, typically         originating from SignM

In both cases, preformatted data is passed into the Translate function by the system. For translation of data destined for Test, the data should be preformatted as per FIG. 370 (all words except the Rs). For translation of signatures for use in WriteAuth, the data should be preformatted as per FIG. 373 (all words except the Rs).

17.2 Input Parameters

Table 265 describes each of the input parameters.

TABLE 265 Description or input parameters for Translate Parameter Description InputKeyRef For translating common key input signature: InputKeyRef.keyNum = Slot number of the key to be used for testing the signature. SIG_(E) produced using K_(InputKeyRef.keyNum) by the external device. InputKeyRef.useChipId = 0 For translating variant key input signatures: InputKeyRef.keyNum = Slot number of the key to be used for generating the variant key. SIG_(E) produced using a variant of K_(InputKeyRef.keyNum) by the external device. InputKeyRef.useChipId = 1 InputKeyRef.chipId = ChipId of the device which generated SIG_(E) using a variant of K_(InputKeyRef.keyNum). DataLength: Length of data in words. Data Data used for testing the input signature and for producing the output signature. R_(E) External random value required for verifying input signature. This will be the R from the input signature generator (i.e device generating SIG_(E)). SIG_(E) External signature required for authenticating input data. The external signature is either generated by a Read function, a Xfer/Rollback function or a Translate function. A correct SIG_(E) = SIG_(KeyRef)(Data | R_(E) | R_(L)). OutputKeyRef For generating common key output signature: OutputKeyRef.keyNum = Slot number of the key for producing the output signature. SIGout produced using K_(OutputKeyRef.keyNum) because the device receiving SIGout shares K_(OutputKeyRef.keyNum) with the translating device. OutputKeyRef.useChipId = 0 For generating variant key output signature: OutputKeyRef.keyNum = Slot number of the key to be used for generating the variant key. SIGout produced using a variant of K_(OutputKeyRef.keyNum) because the device receiving SIGout shares a variant of K_(OutputKeyRef.keyNum) with the translating device. OutputKeyRef.useChipId = 1 OutputKeyRef.chipId = ChipId of the device which receives SIG_(out) produced by a variant of K_(OutputKeyRef.keyNum). R_(E2) External random value required for output signature generation. This will be the R from the destination of SIG_(out). R_(E2) is obtained by calling the Random function on the device which will receive the SIG_(out) from the Translate function. 17.2.1 Input Signature Verification Data Format

This is the same format as used in the Test function. Refer to Section 16.2.1.

17.3 Output Parameters

Table 266 describes each of the output parameters.

TABLE 266 Description of output parameters for Translate Parameter Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1. R_(L2) Local random value used in output signature (i.e SIG_(Out)). SIG_(Out) Output signature produced using OutputKeyRef.keyNum using the data format described in FIG. 372. SIGout = SIG_(OutKeyRef)(Data | R_(L2) | R_(E2)). Refer to Section 10.1.3.1 for details. 17.3.1 SIG_(out)

FIG. 372 shows the data format for output signature generation from the Translate function.

17.4 Function Sequence

The Translate command is illustrated by the following pseudocode:

Accept input parameters-InputKeyRef, DataLength # Accept input parameter- Data based on DataLength For i

0 to (DataLength − 1) Accept next Data EndFor

Accept input parameters—R_(E), SIG_(E), OutputKeyRef, R_(E2)

Check range of InputKeyRef.keyNum and OutputKeyRef.keyNum If invalid ResultFlag

InvalidKey Output ResultFlag Return EndIf #Generate Signature SIG_(L)

GenerateSignature(InputKeyRef,Data,R_(E),R_(L)) # Refer to Figure 371. #Validate input signature If(SIG_(L) = SIG_(E)) Update R_(L) to R_(L2) Else ResultFlag

BadSig Output ResultFlag Return Endif #Generate output signature SIG_(Out)

GenerateSignature(OutputKeyRef,Data,R_(E),R_(L)) # Refer to Figure 372. Update R_(L2) to R_(L3) ResultFlag

Pass Output ResultFlag, R_(L2), SIG_(Out) Return 18 WriteM1+

-   -   Input: VectNum, WordSelect, MVal     -   Output: ResultFlag     -   Changes: M_(VectNum)     -   Availability: All devices         18.1 Function Description

The WriteM1+ function is used to update selected words of M1+, subject to the permissions corresponding to those words stored in P_(VectNum).

Note: Unlike WriteAuth, a signature is not required as an input to this function.

18.2 Input Parameters

Table 267 describes each of the input parameters.

TABLE 267 Description of input parameters for WriteM1+ Parameter Description VectNum Number of the memory vector to be written. Must be in range 1 to (NumVectors − 1) WordSelect Selection of words to be written. 0- indicates corresponding word is not written. 1- indicates corresponding word is to be written as per input. If WordSelect[N bit] is set, then write to M_(VectNum) word N. MVal Multiple of words corresponding to the number of words selected for write. Starts with LSW of M_(VectNum).

Note: Since this function has no accompanying signatures, additional input parameter error checking is required.

18.3 Output Parameters

Table 268 describes each of the output parameters.

TABLE 268 Description of output parameters for WriteM1+ Parameter Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1. 18.4 Function Sequence

The WriteM1+ command is illustrated by the following pseudocode:

Accept input parameters VectNum, WordSelect #Accept MVal as per WordSelect MValTemp[16]

0 # Temporary buffer to hold MVal after being read For i

0 to MaxWordInM # word 0 to word 15 If(WordSelect[i] = 1) Accept next MVal MValTemp[i]

MVal # Store MVal in temporary buffer EndIf EndFor Check range of VectNum If invalid ResultFlag

InvalidVector Output ResultFlag Return EndIf #Checking non authenticated write permission for M1+ PermOK

CheckM1+Perm(VectNum,WordSelect) #Writing M with MVal If(PermOK =1) WriteM(VectNum,MValTemp[ ]) ResultFlag

Pass Else ResultFlag

InvalidPermission EndIf Output ResultFlag Return 18.4.1 PermOK CheckM1+Perm (VectNum, WordSelect)

This function checks WordSelect against permission P_(VectNum) for the selected word

For i

0 to MaxWordInM # word 0 to word 15 If(WordSelect[i] = 1)

(P_(VectNum)[i] = 0) # Trying to write a ReadOnly word Return PermOK

0 EndIf EndFor Return PermOK

 1 18.4.2 WriteM(VectNum, MvalTemp

)

This function copies MvalTemp to M_(VectNum).

For i

0 to MaxWordInM # Copying word from temp buff to M If(VectNum = 1) # If M1 P_(VectNum)[i]

0 # Set permission to ReadOnly before writing EndIf M_(VectNum)[i]

MValTemp[i] # copy word buffer to M word EndIf EndFor 19 WriteFields

-   -   Input FieldSelect, FieldVal     -   Output ResultFlag     -   Changes: M_(VectNum)     -   Availability: All devices         19.1 Function Description

The WriteFields function is used to write new data to selected fields (stored in M0). The write is carried out subject to the non-authenticated write access permissions of the fields as stored in the appropriate words of M1 (see Section 8.1.1.3).

The WriteFields function is used whenever authorization for a write (i.e. a valid signature) is not required. The WriteFieldsAuth function is used to perform authenticated writes to fields. For example, decrementing the amount of ink in an ink cartridge field is permitted by anyone via the WriteFields, but incrementing it during a refill operation is only permitted using WriteFieldsAuth. Therefore WriteFields does not require a signature as one of its inputs.

19.2 Input Parameters

Table 269 describes each of the input parameters.

TABLE 269 Description of input parameters for WriteFields Parameter Description FieldSelect Selection of fields to be written. 0- indicates corresponding field is not written. 1- indicates corresponding field is to be written as per input. If FieldSelect [N bit] is set, then write to Field N of M0. FieldVal Multiple of words corresponding to the words for all selected fields. Since Field0 starts at M0[15], FieldVal words starts with MSW of lower field. Note: Since this function has no accompanying signatures, additional input parameter error checking is required especially if the QA Device communication channel has potential for error. 19.3 Output Parameters

Table 270 describes each of the output parameters.

TABLE 270 Description of output parameters for WriteFields Parameter Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1. 19.4 Function Sequence

The WriteFields command is illustrated by the following pseudocode:

Accept input parameters FieldSelect #Accept FieldVal as per FieldSelect into a temporary buffer MValTemp #Find the size of each FieldNum to accept FieldData FieldSize[16]

0 # Array to hold FieldSize assuming there are 16 fields NumFields

FindNumberOfFieldsInM0(M1,FieldSize) MValTemp[16]

0 # Temporary buffer to hold FieldVal after being read For i

0 to NumFields If FieldSelect[i] = 1 If i = 0 # Check if field number is 0 PreviousFieldEndPos

MaxWordInM Else PreviousFieldEndPos

M1[i−1].EndPos # position of the last word for the # previous field EndIf For j

(PreviousFieldEndPos −1) to M1[FieldNum].EndPos( ) MValTemp[j] = Next FieldVal word #Store FieldVal in MValTemp. EndFor EndIf EndFor #Check non-authenticated write permissions for all fields in FieldSelect PermOK

CheckM0NonAuthPerm(FieldSelect,MValTemp,M0,M1) #Writing M0 with MValTemp if permissions allow writing If(PermOK =1) WriteM(0,MValTemp) ResultFlag

Pass Else ResultFlag

InvalidPermission EndIf Output ResultFlag Return 19.4.1 NumFields FindNumOfFieldsInM0(M1,FieldSize

)

This function returns the number of fields in M0 and an array FieldSize which stores the size of each field.

CurrPos

0 NumFields

0 FieldSize[16]

0 # Array storing field sizes For FieldNum

0 to MaxWordInM If(CurrPos = 0) # check if last field has reached Return FieldNum #FieldNum indicates number of fields in M0 EndIf FieldSizeFieldNum]

CurrPos − M1[FieldNum].EndPos If(FieldSize[FieldNum] < 0) Error # Integrity problem with field attributes Return FieldNum # Lower M0 fields are still valid but higher M0 fields are # ignored Else CurrPos

M1[FieldNum].EndPos EndIf EndFor 19.4.2 Word BitMapForField GetWordMapForField(FieldNum,M1)

This function returns the word bitmap corresponding to a field i.e the field consists of which consecutive words.

WordBitMapForField

0 WordMapTemp

0 PreviousFieldEndPos

M1[FieldNum −1].EndPos # position of the last word for the # previous field For j

(PreviousFieldEndPos +1) to M1[FieldNum].EndPos( ) # Set bit corresponding to the word position WordMapTemp

SHIFTLEFT(1,j) WordBitMapForField

WordMapTemp

WordBitMapForField EndFor Return WordBitMapForField

Return WordBitMapForField

19.4.3 PermOK CheckM0NonAuthPerm(FieldSelect,MValTemp

, M0,M1)

This functions checks non-authenticated write permissions for all fields in FieldSelect.

PermOK CheckM0NonAuthPerm( ) FieldSize[16]

0 NumFields

FindNumOfFieldsInM0(FieldSize) # Loop through all fields in Fieldselect and check their # non-authenticated permission For i

0 to NumFields If FieldSelect[i] = 1 # check selected WordBitMapForField

GetWordMapForField(i,M1) #get word bitmap for field PermOK

CheckAuthFieldPerm(i,WordBitMapForField, MValTemp,M0,) # Check permission for field i in FieldSelect If(PermOK = 0) #Writing is not allowed, return if permissions for field #doesn't allow writing Return PermOK EndIf EndIf EndFor Return PermOK 19.4.4 PermOK

CheckFieldNonAuthPerm(FieldNum,WordBitMapForField, MValTemp

, M0)

This function checks non authenticated write permissions for the field.

DecrementOnly

0 AuthRW

M1[FieldNum].AuthRW NonAuthRW

M1[FieldNum].AuthRW If(NonAuthRW = 0) # No NonAuth write allowed Return PermOK

0 EndIf If((AuthRW = 0)

(NonAuthRW = 1)) # NonAuthRW allowed Return PermOK

1 ElseIf(AuthRW = 1)

(NonAuthRW = 1) # NonAuth DecrementOnly allowed PermOK

CheckInputDataForDecrementOnly(M0,MValTemp, WordBitMapForField) Return PermOK EndIf 19.4.5 PermOK CheckInputDataForDecrementOnly(M0,MValTemp

, WordBitMapForField)

This function checks the data to be written to the field is less than the current value.

DecEncountered

0 LessThanFlag

0 EqualToFlag

0 For i = MaxWordInM to 0 If(WordBitMapForField[i] = 1) # starting word of the field - starting at MSW # comparing the word of temp buffer with M0 current value LessThanFlag

M0[i] < MValTemp[i] EqualToFlag

M0[i] = MValTemp[i] # current value is less or previous value has been decremented If(LessThanFlag =1)

(DecEncountered = 1) DecEncountered

1 PermOK

1 Return PermOK ElseIf(EqualToFlag≠1)  # Only if the value is greater than current and decrement not encountered in previous words PermOK

0 Return PermOK EndIf EndIf EndFor 19.4.6 WriteM(VectNum, MValTemp

)

-   -   Refer to Section 18.4.2 for details.         20 WriteFieldsAuth     -   Input KeyRef, FieldSelect, FieldVal, R_(E), SIG_(E)     -   Output: ResultFlag     -   Changes: _(M0) and R_(L)     -   Availability: All devices         20.1 Function Description

The WriteFieldsAuth command is used to securely update a number of fields (in _(M0)). The write is carried out subject to the authenticated write access permissions of the fields as stored in the appropriate words of M1 (see Section 8.1.1.3). WriteFieldsAuth will either update all of the requested fields or none of them; the write only succeeds when all of the requested fields can be written to.

The WriteFieldsAuth function requires the data to be accompanied by an appropriate signature based on a key that has appropriate write permissions to the field, and the signature must also include the local R (i.e. nonce/challenge) as previously read from this QA Device via the Random function.

The appropriate signature can only be produced by knowing K_(KeyRef). This can be achieved by a call to an appropriate command on a QA Device that holds a key matching K_(KeyRef). Appropriate commands include SignM, XferAmount, XferField, StartXfer, and StartRollBack.

20.2 Input Parameters

Table 271 describes each of the input parameters for WriteAuth.

Parameter Description KeyRef For common key signature generation: KeyRef.keyNum = Slot number of the key to be used for testing the input signature. KeyRef.useChipId = 0 No variant key signature generation required FieldSelect Selection of fields to be written. 0- indicates corresponding field is not written. 1- indicates corresponding field is to be written as per input. If FieldSelect [N bit] is set, then write to Field N of M0. FieldVal Multiple of words corresponding to the total number of words for all selected fields. Since Field0 starts at M0[15], FieldVal words starts with MSW of lower field. RE External random value used to verify input signature. This will be the R from the input signature generator (i.e device generating SIG_(E)). SIGE External signature required for authenticating input data. The external signature is either generated by a Translate or one of the Xfer functions. A correct SIG_(E) = SIG_(KeyRef)(data | R_(E) | R_(L)). 20.2.1 Input Signature Verification Data Format

FIG. 373 shows the input signature verification data format for the WriteAuth function.

-   -   Table 272 gives the parameters included in SIG_(E) for Write         Auth

Value set Length from Parameter in bits Value set internally Input RWSense 3 write constant = 001 Refer to Section 15.3.1.1 FieldNum 4 ● ChipID 48 This QA Device's ChipId FieldData 32 ● per word R_(E) 160 ● R_(L) 160 random value from device 20.3 Output Parameters

Table 273 describes each of the output parameters.

TABLE 273 Description of output parameters for WriteAuth Parameter Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1. 20.4 Function Sequence

The WriteAuth command is illustrated by the following pseudocode:

Accept input parameters-KeyRef, FieldSelect, #Accept FieldVal as per FieldSelect into a temporary buffer MValTemp #Find the size of each FieldNum to accept FieldData FieldSize[16]

0 # Array to hold FieldSize assuming there are 16 fields NumFields

FindNumberOfFieldsInM0(M1,FieldSize) MValTemp[16]

0 # Temporary buffer to hold FieldVal after being read For i

0 to NumFields If i = 0 # Check if field number is 0 PreviousFieldEndPos

MaxWordInM Else PreviousFieldEndPos

M1[i−1].EndPos # position of the last word for the previous field EndIf For j

(PreviousFieldEndPos −1) to M1[FieldNum].EndPos( ) MValTemp[j] = Next FieldVal word #Store FieldVal in MValTemp. EndFor EndIf EndFor Accept R_(E), SIG_(E) Check range of KeyRef.keyNum If invalid range ResultFlag

InvalidKey Output ResultFlag Return EndIf #Generate message for passing to GenerateSignature function data

(RWSense|FieldSelect|ChipId|FieldVal #Generate Signature SIG_(L)

GenerateSignature(KeyRef,data,R_(E),R_(L)) # Refer to Figure 373. #Check signature If(SIG_(L) = SIG_(E)) Update R_(L) to R_(L2) Else ResultFlag

BadSig Output ResultFlag Return EndIf #Check authenticated write permission for all fields in FieldSelect using KeyRef PermOK

CheckM0AuthPerm(FieldSelect,MValTemp,M0,M1,KeyRef) If(PermOK = 1) WriteM(0,MValTemp[ ])# Copy temp buffer to M0 ResultFlag

Pass Else ResultFlag

InvalidPermission EndIf Output ResultFlag Return 20.4.1 PermOK CheckM0AuthPerm(FieldSelect,MValTemp

, M0, M1, KeyRef)

This functions checks non-authenticated write permissions for all fields in FieldSelect using KeyRef.

PermOK CheckM0NonAuthPerm( ) FieldSize[16]

0 NumFields

FindNumOfFieldsInM0(FieldSize) # Loop through fields For i

0 to NumFields If FieldSelect[i] = 1 # check selected WordBitMapForField

GetWordMapForField(i,M1) #get word bitmap for field PermOK

CheckAuthFieldPerm(i,WordBitMapForField,MValTemp,M0, KeyRef)  # Check permission for field i in FieldSelect If(PermOK = 0) #Writing is not allowed, return if #permissions for field doesn't allow writing Return PermOK EndIf EndIf EndFor Return PermOK 20.4.2 PermOK CheckAuthFieldPerm(FieldNum, WordMapForField,MValTemp

, M0,KeyRef)

This function checks authenticated permissions for an M₀ field using KeyRef (whether KeyRef has write permissions to the field).

AuthRW

M1[FieldNum].AuthRW KeyNumAtt

M1[FieldNum].KeyNum If (AuthRW = 0) # Check whether any key has write permissions Return PermOK

0 # No authenticated write permissions EndIf # Check KeyRef has ReadWrite Permission to the field and it is locked If(KeyLock_(KeyNum) = locked)

(KeyNumAtt = KeyRef.keyNum) Return PermOK

1 Else # KeyNum is not a ReadWrite Key KeyPerms

M1[FieldNum].DOForKeys # Isolate KeyPerms for FieldNum # Check Decrement Only Permission for Key If(KeyPerms[KeyRef.keyNum] = 1) # Key is allowed to Decrement field PermOK CheckInputDataForDecrementOnly(M0,MValTemp,WordMapForField) Else # Key is a ReadOnly key PermOK

0 EndIf EndIf Return PermOK 20.4.3 WordBitMapField GetWordMapForField(FieldNum,M1)

-   -   Refer to Section 19.4.2 for details.         20.4.4 PermOK CheckInputDataForDecrementOnly(M0,MValTemp         , WordMapForField)     -   Refer to Section 19.4.5 for details.         20.4.5 WriteM(VectNum, MValTemp         )     -   Refer to Section 18.4.2 for details.         21 SetPerm     -   Input: VectNum, PermVal     -   Output: ResultFlag, NewPerm     -   Changes: P_(n)     -   Availability: All devices         21.1 Function Description

The SetPerm command is used to update the contents of P_(VectNum) (which stores the permission for M_(VectNum)).

The new value for P_(VectNum) is a combination of the old and new permissions in such a way that the more restrictive permission for each part of P_(VectNum) is kept.

M0's permissions are set by M1 therefore they can't be changed.

M1's permissions cannot be changed by SetPerm. M1 is a write-once memory vector and its permissions are set by writing to it.

See Section 8.1.1.3 and Section 8.1.1.5 for more information about permissions.

21.2 Input Parameters

Table 274 describes each of the input parameters for SetPerm.

Parameter Description VectNum Number of the memory vector whose permission is being changed. PermVal Bitmap of permission for the corresponding Memory Vector. Note: Since this function has no accompanying signatures, additional input parameter error checking is required. 21.3 Output Parameters

Table 275 describes each of the output parameters for SetPerm.

Parameter Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1. Perm If VectNum = 0, then no Perm is returned. If VectNum = 1, then old Perm is returned. If VectNum > 1, then new Perm is returned after P_(VectNum) has been changed based on PermVal. 21.4 Function Sequence

The SetPerm command is illustrated by the following pseudocode:

Accept input parameters- Vectum, PermVal Check range of VectNum If invalid ResultFlag

InvalidVector Output ResultFlag Return End If If(VectNum = 0) # No permssions for M0 ResultFlag

Pass Output ResultFlag Return ElseIf(VectNum = 1) ResultFlag

Pass Output ResultFlag Output P₁ Return ElseIf(VectNum >1) # Check that only ‘RW’ parts are being changed # RW(1) → RO(0), RO(0) → RO(0), RW(1) → RW(1) - valid change # RO(0) → RW(1) - Invalid change # checking for change from ReadOnly to ReadWrite temp

~P_(VectNum)

PermVal If(temp = 1 ) # If invalid change is 1 ResultFlag

InvalidPermission Output ResultFlag Else P_(VectNum)

PermVal ResultFlag

Pass Output ResultFlag Output P_(VectNum) EndIf Return EndIf 22 ReplaceKey

-   -   Input: KeyRef, KeyID, KeyLock, EncryptedKey, R_(E), SIG_(E)     -   Output ResultFlag     -   Changes: K_(KeyRef.keyNum) and R_(L)     -   Availability: All devices         22.1 Function Description

The ReplaceKey command is used to replace the contents of a non-locked keyslot, which means replacing the key, its associated keyID, and the lock status bit for the keyslot. A key can only be replaced if the slot has not been locked i.e. the KeyLock for the slot is 0. The procedure for replacing a key also requires knowledge of the value of the current key in the keyslot i.e. you can only replace a key if you know the current key.

Whenever the ReplaceKey function is called, the caller has the ability to make this new key the final key for the slot. This is accomplished by passing in a new value for the KeyLock flag. A new KeyLock flag of 0 keeps the slot unlocked, and permits further replacements. A new KeyLock flag of 1 means the slot is now locked, with the new key as the final key for the slot i.e. no further key replacement is permitted for that slot.

22.2 Input Parameters

Table 276 describes each of the input parameters for Replacekey.

Parameter Description KeyRef For common key signature generation: KeyRef.keyNum = Slot number of the key to be used for testing the input signature, and will be replaced by the new key. KeyRef.useChipId = 0 No variant key signature generation required KeyId KeyId of the new key. The LSB represents whether the new key is a variant or a common key. KeyLock Flag indicating whether the new key should be the final key for the slot or not. (1 = final key, 0 = not final key) EncryptedKey SIG_(Kold)(R_(E)|R_(L)) ⊕ K_(new) where K_(old) = KeyRef.getkey( ). Refer to Section 10.1.3.1 RE External random value required for verifying input signature. This will be the R from the input signature generator (device generating SIG_(E)). In this case the input signature is a generated by calling the GetProgramKey function on a Key Programming device. SIGE External signature required for authenti- cating input data and determining the new key from the EncryptedKey. 22.2.1 Input Signature Generation Data Format

FIG. 374 shows the input signature generation data format for the ReplaceKey function.

Table 277 gives the parameters uincluded in SIG_(E) for ReplaceKey.

Length Value set Value set Parameter in bits internally from Input ChipId 48 This QA Device's ChipId KeyId 32 • R_(E) 160 • EncryptedKey 160 • 22.3 Output Parameters

Table 278 describes each of the output parameters for ReplaceKey.

Parameter Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1. 22.4 Function Sequence

The ReplaceKey command is illustrated by the following pseudocode:

Accept input parameters - KeyRef, KeyId, KeyLock, EncryptedKey,R_(E), SIG_(E) Check KeyRef.keyNum range If invalid ResultFlag

InvalidKey Output ResultFlag Return EndIf #Generate message for passing to GenerateSignature function data

(ChipId|KeyId|KeyLock|R_(E)|EncryptedKey) #Generate Signature SIG_(L)

GenerateSignature(KeyRef,data,Null,Null) # Refer to Figure 374. # Check if the key slot is unlocked If(KeyLock # unlock) ResultFlag

KeyAlreadyLocked Output ResultFlag Return EndIf #Test SIG_(E) If (SIG_(L) # SIG_(E)) ResultFlag

BadSig Output ResultFlag Return EndIf SIG_(L)

GenerateSignature(Key,null,R_(E),R_(L)) Advance R_(L) # Must be atomic - must not be possible to remove power and have KeyId and KeyNum mismatched. Also preferable for KeyLock, although not strictly required. K_(KeyNum)

SIG_(L) ⊕ EncryptedKey KeyId_(KeyNum)

KeyId KeyLock_(KeyNum)

KeyLock ResultFlag

Pass Output ResultFlag Return 23 SignM

-   -   Input: KeyRef, FieldSelect, FieldValLength, FieldVal, ChipID,         R_(E)     -   Output ResultFlag, R_(L), SIG_(out)     -   Changes: R_(L)     -   Availability: Trusted device only         23.1 Function Description

The SignM function is used to generate the appropriate digital signature required for the authenticated write function WriteFieldsAuth. The SignM function is used whenever the caller wants to write a new value to a field that requires key-based write access.

The caller typically passes the new field value as input to the SignM function, together with the nonce (R_(E)) from the QA Device who will receive the generated signature. The SignM function then produces the appropriate signature SIG_(out). Note that SIG_(out) may need to be translated via the Translate function on its way to the final WriteFieldsAuth QA Device.

The SignM function is typically used by the system to update preauthorisation fields (Section 31.4.3).

The key used to produce output signature SIG_(out) depends on whether the trusted device shares a common key or a variant key with the QA Device directly receiving the signature. The KeyRef object passed into the interface must be set appropriately to reflect this.

23.2 Input Parameters

Table 279 describes each of the input parameters for SignM.

Parameter Description KeyRef For generating common key output signature: Ref.keyNum = Slot number of the key for producing the output signature. SIG_(out) produced using K_(KeyRef.keyNum) because the device receiving SIG_(out) shares K_(KeyRef.keyNum) with the trusted device. KeyRef.useChipId = 0 For generating variant key output signature: KeyRef.keyNum = Slot number of the key to be used for generating the variant key. SIG_(out) produced using a variant of K_(KeyRef.keyNum) because the device receiving SIG_(out) shares a variant of K_(KeyRef.keyNum) with the trusted device. KeyRef.useChipId = 1 KeyRef.chipId = ChipId of the device which receives SIG_(out). FieldNum Field number of the field that will be written to. FieldDataLength The length of the FieldData in words. FieldData The value that will be written to the field selected by FieldNum. R_(E) External random value used in the output signature generation. R_(E) is obtained by calling the Random function on the device, which will receive the SIG_(out) from the SignM function, which in this case is the WriteAuth function or the Translate function. ChipId Chip identifier of the device whose WriteAuth function will be called subsequently to perform an authenticated write to its FieldNum of M0. 23.3 Output Parameters

Table 280 describes each of the output parameters.

TABLE 280 Description of output parameters for SignM Parameter Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1. R_(L) Internal random value used in the output signature. SIG_(out) SIG_(Out) = SIG_(KeyRef)(data | R_(L) | R_(E)) as shown in FIG. 373. As per FIG. 373, R_(E) is actually R_(L) and R_(L)is R_(E) with respect to device producing SIG_(out) to be applied to WriteAuth function. 23.3.1 SIG_(out)

-   -   Refer to Section 20.2.1.         23.4 Function Sequence

The SignM command is illustrated by the following pseudocode:

Accept input parameters - KeyRef, FieldNum, FieldDataLength # Accept FieldData words For i = 0 to FieldValLength Accept next FieldData EndFor Accept ChipId, R_(E) Check KeyRef.keyNum range If invalid ResultFlag

InvalidKey Output ResultFlag Return EndIf #Generate message for passing into the GenerateSignature function data

(RWSense|FieldSelect|ChipId|FieldVal) #Generate Signature SIG_(out)

GenerateSignature(KeyRef,data,R_(L),R_(E)) # Refer to Section 20.2.1. Advance R_(L)to R_(L2) ResultFlag

Pass Output parameters ResultFlag, R_(L,)SIG_(out) Return Functions On A key programming qa device 24 Concepts

The key programming device is used to replace keys in other devices.

The key programming device stores both the old key which will be replaced in the device being programmed, and the new key which will replace the old key in the device being programmed. The keys reside in normal key slots of the key programming device.

Any key stored in the key programming device can be used as an old key or a new key for the device being programmed, provided it is permitted by the key replacement map stored within the key programming device.

FIG. 375 is representation of a key replacement map. The 1 s indicates that the new key is permitted to replace the old key. The 0s indicates that key replacement is not permitted for those positions. The positions in FIG. 13 which are blank indicate a 0.

According to the key replacement map in FIG. 13, K₅ can replace K₁, K₆ can replace K₃, K₄, K₅,K₇, K₃ can replace K₂, K₀ can replace K₂, and K₂ can replace K₆. No key can replace itself.

FIG. 375._Key Replacement Map

The key replacement map must be readable from an external system and must be updateable by an authenticated write. Therefore, the key replacement map must be stored in an M0 field. This requires one of the keys residing in the key programming device to be have ReadWrite access to the key replacement map. This key is referred to as the key replacement map key and is used to update the key replacement map.

There will one key replacement map field in a key programming device.

No key replacement mappings are allowed to the key replacement map key because it should not be used in another device being programmed. To prevent the key replacement map key from being used in key replacement, in case the mapping has been accidentally changed, the key replacement map key is allocated a fixed key slot of 0 in all key programming devices. If a GetProgram function is invoked on the key programming device with the key replacement map key slot number 0 it immediately returns an error, even before the key replacement map is checked.

The keys K₀ to K₇ in the key programming device are initially set during the instantiation of the key programming device. Thereafter, any key can be replaced on the key programming device by another key programming device If a key in a key slot of the key programming device is being replaced, the key replacement map for the old key must be invalidated automatically. This is done by setting the row and column for the corresponding key slot to 0 For example, if K₁ is replaced, then column 1 and row 1 are set to 0, as indicated in FIG. 376.

The new mapping information for K, is then entered by performing an authenticated write of the key replacement map field using the key replacement map key.

24.1 Key Replacement Map Data Structure

As mentioned in Section 24, the key replacement map must be readable by external systems and must be updateable using an authenticated write by the key replacement map key. Therefore, the key replacement map is stored in an M0 field of the key programming device. The map is 8×8 bits in size and therefore can be stored in a two word field. The LSW of key replacement map stores the mappings for K₀-K₃. The MSW of key replacement map stores the mappings for K₄-K₇. Referring to FIG. 375, key replacement map LSW is 0x40092000 and MSW is 0x40224040. Referring to FIG. 376, after K₁ is replaced in the key programming device, the value of the key replacement map LSW is 0x40090000 and MSW is 0x40224040.

The key replacement map field has an M1 word representing its attributes. The attribute setting for this field is specified in Table 281.

TABLE 281 Key replacement map attribute setting Attribute name Value Explanation Type TYPE_KEY_MAP Indicates that the field Refer to Appendix A. value represents a key re- placement map. Only one such field per key programming QA Device. KeyNum 0 Slot number of the key replacement map key. NonAuthRW 0 No non authenticated writes is permitted. AuthRW 1 Authenticated write is permitted. KeyPerms 0 No Decrement Only permission for any key. EndPos Value such that field size is 2 words 24.2 Basic Scheme

The Key Replacement sequence is shown FIG. 377.

Following is a sequential description of the transfer and rollback process:

-   1. The System gets a Random number from the QA Device whose keys are     going to be replaced. -   2. The System makes a GetProgramKey Request to the Key Programming     QA Device. The Key Programming QA Device must contain both keys for     QA Device whose keys are being replaced—Old Keys which are the keys     that exist currently (before key replacement), and the New Keys     which are the keys which the QA Device will have after a successful     processing of the ReaplceKey Request. The GetProgramKey Request is     called with the Key number of the Old Key (in the Key Programming QA     Device) and the Key Number of the New Key (in the Key Programming QA     Device), and the Random number from (1). The Key Programming QA     Device validates the GetProgramKey Request based on the     KeyReplacement map, and then produces the necessary GetProgramKey     Output. The GetProgramKey Output consists of the encrypted New Key     (encryption done using the Old Key), along with a signature using     the Old Key. -   3. The System then applies GetProgramKey Output to the QA Device     whose key is being replaced, by calling the ReplaceKey function on     it, passing in the GetProgramKey Output. The ReplaceKey function     will decrypt the encrypted New Key using the Old Key, and then     replace its Old Key with the decrypted New Key.     25 Functions     25.1 GetProgamKey     -   Input: OldKeyRef, ChipID, R_(E), KeyLock, NewKeyRef     -   Output: ResultFlag, R_(L), EncryptedKey, KeyIdOfNewKey,         SIGo_(out)     -   Changes: R_(L)     -   Availability: Key programming device         25.1.1 Function Description

The GetProgramKey works in conjunction with the ReplaceKey command, and is used to replace the specified key and its KeyId. This function is available on a key programming device and produces the necessary inputs for the ReplaceKey function. The ReplaceKey command is then run on the device whose key is being replaced.

The key programming device must have both the old key and the new key programmed as its keys, and the key replacement map stored in one of its M0 field, before GetProgramKey can be called on the device.

Depending on the OldKeyRef object and the NewKeyRef object passed in, the GetProgramKey will produce a signature to replace a common key by a common key, a variant key by a common key, a common key by a variant key or a variant key by a variant key.

25.1.2 Input Parameters

Table 282 describes each of the input parameters for GetProgramKey.

Parameter Description OldKeyRef Old key is a common key: OldKeyRef.keyNum = Slot number of the old key in the Key Programming QA Device. The device whose key is being replaced, shares a common key K_(OldKeyRef.keyNum) with the key programming device. OldKeyRef.useChipId = 0 Old key is a variant key KeyRef.keyNum = Slot number of the old key in the Key Programming QA Device. that will be used to generate the variant key. The device whose key is being replaced, shares a variant of K_(OldKeyRef.keyNum) with the key programming device. OldKeyRef.useChipId = 1 OldKeyRef.chipId = ChipId of the device whose variant of K_(OldKeyRef.keyNum) key is being replaced. ChipId Chip identifier of the device whose key is being replaced. RE External random value which will be used in output signature generation. R_(E) is obtained by calling the Random function on the device being programmed. This will also receive the SIGout from the GetProgramKey function. SIGout is passed in to ReplaceKey function. KeyLock Flag indicating whether the new key should be unlocked/locked into its slot. NewKeyRef New key is a common key: NewKeyRef.keyNum = Slot number of the new key in the Key Programming QA Device. The device whose key is being replaced, will receive a common key K _(NewKeyRef.keyNum) from the key programming device. NewKeyRef.useChipId = 0 NewKey is a variant key: NewKeyRef.keyNum = Slot number of the new key in the Key Programming QA Device. that will be used to generate the new variant key. The device whose key is being replaced, will receive a new key which is a variant of K_(NewKeyRef.keyNum)from the key programming device. NewKeyRef.useChipId = 1 NewKeyRef.chipId = ChipId of the device receiving a new key, the new key is a variant of the K_(NewKeyRef.keyNum). 25.1.3 Output Parameters

Table 283 describes each of the output parameters for GetProgramKey.

Parameter Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1 and Table 284 R_(L) Internal random value used in the output signature. EncryptedKey SIG_(Kold)(R_(L)|R_(E)) ⊕ K_(new) KeyIdOfNewKey KeyId of the new key. The LSB represents whether the new key is a variant or a common key. SIG_(out) SIG_(out) = SIG_(Kold)(data | R_(L) | R_(E))

TABLE 284 ResultFlag definitions for GetProgramKey Result Flag Description InvalidKeyReplacementMap Key replacement map field invalid or doesn't exist. KeyReplacementNotAllowed Key replacement not allowed as per key replacement map. 25.1.3.1 SIG_(out)

-   -   FIG. 378 shows the output signature generation data format for         the GetProgramKey function.         25.1.4 Function Sequence     -   The GetProgramKey command is illustrated by the following         pseudocode:

Accept input parameters - OldKeyRef, ChipId, R_(E), KeyLock, NewKeyRef ---------------------------------------------------------------- # key replacement map key stored in K0, must not be used for key replacement If(OldKeyRef.keyNum = 0)

(NewKeyRef.keyNum = 0) ResultFlag

Fail Output ResultFlag Return EndIf ---------------------------------------------------------------- CheckRange(OldKeyRef.keyNum) If invalid ResultFlag

InvalidKey Output ResultFlag Return EndIf ---------------------------------------------------------------- CheckRange(NewKeyRef.keyNum) If invalid ResultFlag

InvalidKey Output ResultFlag Return EndIf ---------------------------------------------------------------- # Find M0 words that represent the key replacement map WordSelectForKeyMapField

GetWordSelectForKeyMapField(M1) If(WordSelectForKeyMapField = 0) ResultFlag

InvalidKeyReplacementMap Output ResultFlag Return EndIf ---------------------------------------------------------------- #CheckMapPermits key replacement ReplaceOK

CheckMapPermits(WordSelectForKeyMapField,OldKeyNum, NewKeyNum) If(ReplaceOK = 0) ResultFlag

KeyReplacementNotAllowed Output ResultFlag Return EndIf ---------------------------------------------------------------- #All checks are OK, now generate Signature with OldKey SIG_(L)

GenerateSignature(OldKeyRef,null,R_(L),R_(E)) #Get new key K_(NewKey)

NewKeyRef.getKey( ) #Generate Encrypted Key EncryptedKey

SIG_(L) ⊕ K_(NewKey) #Set base key or variant key - bit 0 of KeyId If(NewKeyRef.useChipId = 1) KeyId

0x0001

0x0001 Else KeyId

0x0001

0x0000 EndIf #Set the new key KeyId to the KeyId - bits 1–30 of KeyId KeyIdOfNewKey

SHIFTLEFT(KeyIdOfNewKey,1) KeyId

KeyId

KeyIdOfNewKey #Set the KeyLock as per input - bit 31 of KeyId KeyLock

SHIFTLEFT(KeyLock,31) #KeyId

KeyId

KeyLock #Generate message for passing in to the GenerateSignature function data

ChipId|KeyId|R_(L)|EncryptedKey #Generate output signature SIG_(out)

GenerateSignature(OldKeyRef,data,null,null) # Refer to Figure 378 Advance R_(L)to R_(L2) ResultFlag

Pass Output ResultFlag, R_(L,)SIG_(out),KeyId, EncryptedKey Return 25.1.4.1 WordSelectForField GetWordSelectForKeyMapField(M1)

This function gets the words corresponding to the key replacement map in M0.

FieldSize[16]

0 # Array to hold FieldSize assuming there are 16 fields NumFields

FindNumberOfFieldsInM0(M1,FieldSize) #Find the key replacement map field For i

0 to NumFields If(TYPE_KEY_MAP = M1[i].Type) # Field is key map field MapFieldNum

i Return Endif EndFor #Get the words corresponding to the key replacement map WordMapForField

GetWordMapForField(MapFieldNum,M1) Return WordSelectForField 25.1.4.2 NumFields FindNumOfFieldsInM0(M1, FieldSize

)

-   -   Refer to Figure 19.4.1 for details         25.1.4.3 WordMapForField GetWordMapForField(FieldNum,M1)     -   Refer to Section 19.4.2 for details.         25.1.4.4 ReplaceOK CheckMapPermits(WordSelectForKeyMapField,         OldKeyNum, NewKeyNum,M0)     -   This function checks whether key replacement map permits key         replacement.     -   #Isolate KeyReplacementMap based on WordSelectForKeyMapField and         M0 KeyReplacementMap[64 bit]     -   #Isolate permission bit corresponding for NewKeyNum in the map         for OldKeyNm     -   ReplaceOK←KeyReplacementMap[(OldKeyNum×8+NewKeyNum) bit]     -   Return ReplaceOK         25.2 ReplaceKey     -   Input KeyRef, KeyID, KeyLock, EncryptedKey,R_(E), SIG_(E)     -   Output: ResultFlag     -   Changes: K_(KeyNum) and R_(L)     -   Availability: Key programming device         25.2.1 Function Description

This function is used for replacing a key in a key programming device and is similar to the generic ReplaceKey function(Refer to Section 24), with an additional step of setting the KeyRef.keyNum column and KeyRef.keyNum row key replacement map to 0.

25.2.2 Input Parameters

Refer to Section 22.

25.2.3 Output Parameters

Refer to Section 22.

25.2.4 Function Sequence

The ReplaceKey command is illustrated by the following pseudocode:

Accept input parameters - KeyRef, KeyId, EncryptedKey,R_(E), SIG_(E) #Generate message for passing into GenerateSignature function data

(ChipId|KeyId|R_(E)|EncryptedKey)# Refer to Figure 374. ---------------------------------------------------------------- # Validate KeyRef, and then verify signature ResultFlag = ValidateKeyRefAndSignature(KeyRef,data,R_(E),R_(L)) If (ResultFlag ≠ Pass) Output ResultFlag Return EndIf ---------------------------------------------------------------- # Check if the key slot is unlocked Isolate KeyLock for KeyRef If(KeyLock = lock) ResultFlag

KeyAlreadyLocked Output ResultFlag Return EndIf SIG_(L)

GenerateSignature(Key,Null,R_(E),R_(L)) Advance R_(L) # Find M0 words that represent the key replacement map WordSelectForKeyMapField

GetWordSelectForKeyMapField(M1) # Set the bits corresponding to the KeyRef.keyNum row and column to 0 # i.e invalidate the key replacement map for KeyRef.keyNum. #Must be done before the key is replaced and must be atomic with key replacement. SetFlag

SetKeyMapForKeyNum(WordSelectForKeyMapField, KeyRef.keyNum,M0) If(SetFlag = 1) # Must be atomic - must not be possible to remove power and have KeyId and KeyNum mismatched K_(KeyNum)

SIG_(L) ⊕ EncryptedKey KeyId_(KeyNum)

KeyId KeyLock_(KeyNum)

KeyLock ResultFlag

Pass Else ResultFlag

Fail EndIf Output ResultFlag Return 25.2.4.1 WordSelectForField GetWordSelectForKeyMapField(M1)

Refer to Figure 25.1.4.1 for details.

25.2.4.2 SetFlag SetKeyMapForKeyNum(WordSelectForKeyMapField,KeyNum, M0)

This function invalidates the key replacement map for KeyNum.

#Isolate KeyReplacementMap based on WordSelectForKeyMapField and M0 KeyReplacementMap[64 bit] # Set KeyNum row (all bits) to 0 in the KeyReplacementMap For i = 0 to 7 KeyReplacementMap[(KeyNum × 8 + i)bit]

0 EndFor # Set KeyNum column to 0 in the KeyReplacementMap For i = 0 to 7 KeyReplacementMap[(i×8 + KeyNum)bit]

0 EndFor SetFlag

1 Return SetFlag Functions Upgrade Device (Ink Re/fill) 26 Concepts 26.1 Purpose

In a printing application, an ink cartridge contains an Ink QA Device storing the ink-remaining values for that ink cartridge. The ink-remaining values decrement as the ink cartridge is used to print.

When an ink cartridge is physically re/filled, the Ink QA Device needs to be logically re/filled as well. Therefore, the main purpose of an upgrade is to re/fill the ink-remaining values of an Ink QA Device in an authorised manner.

The authorisation for a re/fill is achieved by using a Value Upgrader QA Device which contains all the necessary functions to re/write to the Ink QA Device. In this case, the value upgrader is called an Ink Refill QA Device, which is used to fill/refill ink amount in an Ink QA Device.

When an Ink Refill QA Device increases (additive) the amount of ink-remaining in an Ink QA Device, the amount of ink-remaining in the Ink Refill QA Device is correspondingly decreased. This means that the Ink Refill QA Device can only pass on whatever ink-remaining value it itself has been issued with. Thus an Ink Refill QA Device can itself be replenished or topped up by another Ink Refill QA Device.

The Ink Refill QA Device can also be referred to as the Upgrading QA Device, and the Ink QA Device can also be referred to as the QA Device being upgraded.

The refill of ink can also be referred to as a transfer of ink, or transfer of amount/valu, or an upgrade.

Typically, the logical transfer of ink is done only after a physical transfer of ink is successful.

26.2 Requirements

The transfer process has two basic requirements:

-   -   The transfer can only be performed if the transfer request is         valid. The validity of the transfer request must be completely         checked by the Ink Refill QA Device, before it produces the         required output for the transfer. It must not be possible to         apply the transfer output to the Ink QA Device, if the Ink         Refill QA Device has been already been rolled back for that         particular transfer.     -   A process of rollback is available if the transfer was not         received by the Ink QA Device. A rollback is performed only if         the rollback request is valid. The validity of the rollback         request must be completely checked by the Ink Refill QA Device,         before it adjusts its value to a previous value before the         transfer request was issued. It must not be possible to rollback         an Ink Refill QA Device for a transfer which has already been         applied to the Ink QA Device i.e the Ink Refill QA Device must         only be rolled back for transfers that have actually failed.         26.3 Basic Scheme

The transfer and rollback process is shown in FIG. 379.

Following is a sequential description of the transfer and rollback process:

-   1. The System Reads the memory vectors M0 and M1 of the Ink QA     Device. The output from the read which includes the M0 and M1 words     of the Ink QA Device, and a signature, is passed as an input to the     Transfer Request. It is essential that M0 and M1 are read together.     This ensures that the field information for M0 fields are correct,     and have not been modified, or substituted from another device.     Entire M0 and M1 must be read to verify the correctness of the     subsequent Transfer Request by the Ink Refill QA Device. -   2. The System makes a Transfer Request to the Ink Refill QA Device     with the amount that must be transferred, the field in the Ink     Refill QA Device the amount must be transferred from, and the field     in Ink QA Device the amount must be transferred to. The Transfer     Request also includes the output from Read of the Ink QA Device. The     Ink Refill QA Device validates the Transfer Request based on the     Read output, checks that it has enough value for a successful     transfer, and then produces the necessary Transfer Output. The     Transfer Output typically consists of new field data for the field     being refilled or upgraded, additional field data required to ensure     the correctness of the transfer/rollback, along with a signature. -   3. The System then applies the Transfer Output to the Ink QA Device,     by calling an authenticated Write function on it, passing in the     Transfer Output. The Write is either successful or not. If the Write     is not successful, then the System will repeat calling the Write     function using the same transfer output, which may be successful or     not. If unsuccesful the System will initiate a rollback of the     transfer. The rollback must be performed on the Ink Refill QA     Device, so that it can adjust its value to a previous value before     the current Transfer Request was initiated. It is not necessary to     perform a rollback immediately after a failed Transfer. The Ink QA     Device can still be used to print, if there is any ink remaining in     it. -   4. The System starts a rollback by Reading the memory vectors M0 and     M1 of the Ink QA Device. -   5. The System makes a StartRollBack Request to the Ink Refill QA     Device with same input parameters as the Transfer Request, and the     output from Read in (4). The Ink Refill QA Device validates the     StartRollBack Request based on the Read output, and then produces     the necessary Pre-rollback output. The Pre-rollback output consists     only of additional field data along with a signature. -   6. The System then applies the Pre-rollback Output to the Ink QA     Device, by calling an authenticated Write function on it, passing in     the Pre-rollback output. The Write is either successful or not. If     the Write is not successful, then either (6), or (5) and (6) must be     repeated. -   7. The System then Reads the memory vectors M0 and M1 of the Ink QA     Device. -   8. The System makes a RollBack Request to the Ink Refill QA Device     with same input parameters as the Transfer Request, and the output     from Read (7). The Ink Refill QA Device validates the RollBack     Request based on the Read output, and then rolls back its field     corresponding to the transfer.     26.3.1 Transfer

As we mentioned, the Ink QA Device stores ink-remaining values in its M0 fields, and its corresponding M₁ words contains field information for its ink-remaining fields. The field information consists of the size of the field, the type of data stored in field and the access permission to the field. See Section 8.1.1 for details.

The Ink Refill QA Device also stores its ink-remaining values in its M0 fields, and its coressponding M₁ words contains field information for its ink-remaining fields.

26.3.1.1 Authorisation

The basic authorisation for a transfer comes from a key, which has authenticated ReadWrite permission (stored in field information as KeyNum) to the ink-remaining field (to which ink will be transferrred) in the Ink QA Device. We will refer to this key as the refill key. The refill key must also have authenticated decrement-only permission for the ink-remaining field (from which ink will be transferred) in the Ink Refill QA Device.

After validating the input transfer request, the Ink Refill QA Device will decrement the amount to be transferred from its ink-remaining field, and produce a transfer amount (previous ink-remaining amount in the Ink QA Device+transfer amount), additional field data, and a signature using the refill key. Note that the Ink Refill QA Device can decrement its ink-remaining field only if the refill key has the permission to decrement it.

The signature produced by the Ink Refill QA Device is subsequently applied to the Ink QA Device. The Ink QA Device will accept the transfer amount only if the signature is valid. Note that the signature will only be valid if it was produced using the refill key which has write permission to the ink-remaining field being written.

26.3.1.2 Data Type Matching

The Ink Refill QA Device validates the transfer request by matching the Type of the data in ink-remaining information field of Ink QA Device to the Type of data in ink-remaining information field of the Ink Refill QA Device. This ensures that equivalent data Types are transferred i.e Network_OEM1_infrared ink is not transferred to Network_OEM1_cyan ink.

26.3.1.3 Addition Validation

Additional validation of the transfer request must also be performed before a transfer output is generated by the Ink Refill QA Device. These are as follows:

-   -   For the Ink Refill QA Device:

-   1. Whether the field being upgraded is actually present.

-   2. Whether the field being upgraded can hold the upgraded amount.     -   For the Ink QA Device:

-   1. Whether the field from which the amount is transferred is     actually present.

-   2. Whether the field has sufficient amount required for the     transfer.     26.3.1.4 Rollback Facilitation

To facilitate a rollback, the Ink Refill QA Device will store a list of transfer requests processed by it. This list is referred to as the Xfer Entry cache. Each record in the list consists of the transfer parameters corresponding to the transfer request.

26.3.2 Rollback

A rollback request is validated by looking through the Xfer Entry of the Ink Refill QA Device and finding the request that should be rolled back. After the right transfer request is found the Ink Refill QA Device checks that the output from the transfer request was not applied to the Ink QA Device by comparing the current Read of the Ink QA Device to the values in the Xfer Entry cache, and finally rolls back its ink-remaining field (from which the ink was transferred) to a previous value before the transfer request was issued.

The Ink Refill QA Device must be absolutely sure that the Ink QA Device didn't receive the transfer. This factor determines the additional fields that must be written along with transfer amount, and also the parameters of the transfer request that must be stored in the Xfer Entry cache to facilitate a rollback, to prove that the Printer QA Device didn't actually receive the transfer.

26.3.2.1 Sequence Fields

The rollback process must ensure that the transfer output (which was previously produced) for which the rollback is being performed, cannot be applied after the rollback has been performed. How do we achieve this? There are two separate decrement-only sequence fields (SEQ_(—)1 and SEQ_(—)2) in the Ink QA Device which can only be decremented by the Ink Refill QA Device using the refill key. The nature of data to be written to the sequence fields is such that either the transfer output or the pre-rollback output can be applied to the Ink QA Device, but not both i.e they must be mutually exclusive.Refer to Table 285 for details.

TABLE 285 Sequence field data for Transfer and Pre-rollback Sequence Field data written to Ink QA Device Function SEQ_1 SEQ_2 Explanation Initialised 0xFFFFFFFF 0xFFFFFFFF Written using the sequence key which is different from the refill key Write using (Previous Value − 2) (Previous Value − 1) Written using the refill key using the Transfer If Previous Value = If Previous Value = refill key which has decrement-only Output intialised value intialised value permission on the fields. Value cannot be then 0xFFFFFFFD then 0xFFFFFFFE written if pre-rollback output is already written. Write using (Previous Value − 1) (Previous Value − 2) Written using the refill key using the Pre-rollback If Previous Value = If Previous Value = refill key which has decrement-only intialised value initialised value permission on the fields. Value can be then 0xFFFFFFFE then 0xFFFFFFFD written only if Transfer Output has not been written.

The two sequence fields are initialised to 1xFFFFFFFF using sequence key. The sequence key is different to the refill key, and has authenticated ReadWrite permission to both the sequence fields. The transfer output consists of the new data for the field being upgraded, field data of the two sequence fields, and a signature using the refill key. The field data for SEQ_(—)1 is decremented by 2 from the original value that was passed in with the transfer request. The field data for SEQ_(—)2 is decremented by 1 from the original value that was passed in with the transfer request.

The pre-rollback output consists only of the field data of the two sequence fields, and a signature using the refill key. The field data for SEQ_(—)1 is decremented by 1 from the original value that was passed in with the transfer request. The field data for SEQ_(—)2 is decremented by 2 from the original value that was passed in with the transfer request.

Since the two sequence fields are decrement-only fields, the writing of the transfer output to QA Device being upgraded will prevent the writing of the pre-rollback output to QA Device being upgraded. If the writing of the transfer output fails, then pre-rollback can be written. However, the transfer output cannot be written after the pre-rollback has been written.

Before a rollback is performed, the Ink Refill QA Device must confirm that the sequence fields was successfully written to the pre-rollback values in the Ink QA Device. Because the sequence fields are Decrement-Only fields, the Ink QA Device will allow pre-rollback output to be written only if the upgrade output has not been written. It also means that the transfer output cannot be written after the pre-rollback values have been written.

26.3.2.1.1 Field Information of the Sequence Data Field

For a device to be upgradeable the device must have two sequence fields SEQ_(—)1 and SEQ_(—)2 which are written with sequence data during the transfer sequence. Thus all upgrading QA devices, ink QA Devices and printer QA Devices must have two sequence fields. The upgrading QA Devices must also have these fields because they can be upgraded as well.

The sequence field information is defined in Table 286.

TABLE 286 Sequence field information Attribute Name Value Explanation Type TYPE_SEQ_1 or See Appendix A for TYPE_SEQ_2. exact value. KeyNum Slot number of Only the sequence key the sequence key. has authenticated ReadWrite access to this field. Non Auth 0 Non authenticated RW Perm ReadWrite is not allowed to the field. Auth RW 1 Authenticated (key Perm based) ReadWrite access is allowed to the field. KeyPerm KeyPerms[KeyNum] = 0 KeyNum is the slot number of the sequence key, which has ReadWrite permission to the field. KeyPerms[Slot Refill key can number of the refill decrement the key] = 1 sequence field. KeyPerms[others = All other keys have 0 . . . 7(except ReadOnly access. refill key)] = 0 End Pos Set as required. Size is typically 1 word. 26.3.3 Upgrade States

There are three states in an transfer sequence, the first state is initiated for every transfer, while the next two states are initiated only when the transfer fails. The states are—Xfer, StartRollback, and Rollback.

26.3.3.1 Upgrade Flow

FIG. 380 shows a typical upgrade flow.

26.3.3.2 Xfer

This state indicates the start of the transfer process, and is the only state required if the transfer is successful. During this state, the Ink Refill QA Device adds a new record to its Xfer Entry cache, decrements its amount, produces new amount, new sequence data (as described in Section 26.3.2.1) and a signature based on the refill key.

The Ink QA Device will subsequently write the new amount and new sequence data, after verifying the signature. If the new amount can be successfully written to the Ink QA Device, then this will finish a successful transfer.

If the writing of the new amount is unsuccessful (result returned is BAD SIG), the System will re-transmit the transfer output to the Ink QA Device, by calling the authenticated Write function on it again, using the same transfer output.

If retrying to write the same transfer output fails repeatedly, the System will start the rollback process on Ink Refill QA Device, by calling the Read function on the Ink QA Device, and subsequently calling the StartRollBack function on the Ink Refill QA Device. After a successful rollback is performed, the System will invoke the transfer sequence again.

26.3.3.3 StartRollBack

This state indicates the start of the rollback process. During this state, the Ink Refill QA Device produces the next sequence data and a signature based on the refill key. This is also called a pre-rollback, as described in Section 26.3.2.

The pre-rollback output can only be written to the Ink QA Device, if the previous transfer output has not been written. The writing of the pre-rollback sequence data also ensures, that if the previous transfer output was captured and not applied, then it cannot be applied to the Ink QA Device in the future.

If the writing of the pre-rollback output is unsuccessful (result returned is BAD SIG), the System will re-transmit the pre-rollback output to the Ink QA Device, by calling the authenticated Write function on it again, using the same pre-rollback output.

If retrying to write the same pre-rollback output fails repeatedly, the System will call the StartRollback on the Ink Refill QA Device again, and subsequently calling the authenticated Write function on the Ink QA Device using this output.

26.3.3.4 Rollback

This state indicates a successful deletion (completion) of a transfer sequence. During this state, the Ink Refill QA Device verifies the sequence data produced from StartRollBack has been correctly written to Ink Refill QA Device, then rolls its ink-remaining field to a previous value before the transfer request was issued.

26.3.4 Xfer Entry Cache

The Xfer Entry data structure must allow for the following:

-   -   Stores the transfer state and sequence data for a given transfer         sequence.     -   Store all data corresponding to a given transfer, to facilitate         a rollback to the previous value before the transfer output was         generated.

The Xfer Entry cache depth will depend on the QA Chip Logical Interface implementation. For some implementations a single Xfer Entry value will be saved. If the Ink Refill QA Device has no powersafe storage of Xfer Entry cache, a power down will cause the erasure of the Xfer Entry cache and the Ink Refill QA Device will not be able to rollback to a pre-power-down value.

A dataset in the Xfer Entry cache will consist of the following:

-   -   Information about the QA Device being upgraded:

-   a. ChipId of the device.

-   b. FieldNum of the M0 field (i.e what was being upgraded).     -   Information about the upgrading QA Device:

-   a. FieldNum of the M0 field used to transfer the amount from.     -   XferVal—the transfer amount.     -   Xfer State—indicating at which state the transfer sequence is.         This will consist of:

-   a. State definition which could be one of the following: —Xfer,     StartRollBack and complete/deleted.

-   b. The value of sequence data fields SEQ_(—)1 and SEQ_(—)2.     26.3.4.1 Adding New Dataset

A new dataset is added to Xfer Entry cache by the Xfer function.

There are three methods which can be used to add new dataset to the Xfer Entry cache. The methods have been listed below in the order of their priority:

-   1. Replacing existing dataset in Xfer Entry cache with new dataset     based on ChipId and FieldNum of the Ink QA Device in the new     dataset. A matching ChipId and FieldNum could be found because a     previous transfer output corresponding to the dataset stored in the     Xfer Entry cache has been correctly received and processed by the     Ink Refill QA Device, and a new transfer request for the same Ink QA     Device, same field, has come through to the Ink Refill QA Device. -   2. Replace existing dataset cache with new dataset based on the Xfer     State. If the Xfer State for a dataset indicates deleted (complete),     then such a dataset will not be used for any further functions, and     can be overwritten by a new dataset. -   3. Add new dataset to the end of the cache. This will automatically     delete the oldest dataset from the cache regardless of the Xfer     State.     26.4 Different Types of Transfer

There can be three types of transfer:

-   -   Peer to Peer Transfer—This transfer could be one of the 2 types         described below:

-   a. From an Ink Refill QA Device to a Ink QA Device. This is     performed when the Ink QA Device is refilled by the Ink Refill QA     Device.

-   b. From one Ink Refill QA Device to another Ink Refill QA Device,     where both QA Devices belong to the same OEM. This is typically     performed when OEM divides ink from one Ink Refill QA Device to     another Ink Refill QA Device, where both devices belong to the same     OEM     -   Heirachical Transfer—This is a transfer from one Ink Refill QA         Device to another Ink Refill QA Device, where the QA Devices         belong to different organisation, say ComCo and OEM. This is         typically performed when ComCo divides ink from its refill         device to several refill devices belonging to several OEMs.

FIG. 381 is a representation of various authorised ink refill paths in the printing system.

26.4.1 Hierarchical Transfer

Referring to FIG. 381, this transfer is typically performed when ink is transferred from ComCo's Ink Refill QA Device to OEM's Ink Refill QA Device, or from QACo's Ink Refill QA Device to ComCo's Ink Refill QA Device.

26.4.1.1 Keys and Access Permission

We will explain this using a transfer from ComCo to OEM.

There is an ink-remaining field associated with the ComCo's Ink Refill QA Device. This ink-remaining field has two keys associated with:

-   -   The first key transfers ink to the device from another refill         device (which is higher in the heirachy), fills/refills         (upgrades) the device itself. This key has authenticated         ReadWrite permission to the field.     -   The second key transfers ink from it to other devices (which are         lower in the heirachy), fills/refills (upgrades) other devices         from it. This key has authenticated decrement-only permission to         the field.

There is an ink-remaining field associated with the OEM's Ink refill device. This ink-remaining field has a single key associated with:

-   -   This key transfers ink to the device from another refill device         (which is higher or at the same level in the hierarchy),         fills/refills (upgrades) the device itself, and additionally         transfers ink from it to other devices (which are lower in the         heirachy), fills/refills (upgrades) other devices from it.         Therefore, this key has both authenticated ReadWrite and         decrement-only permission to the field. For a successful         transfer ink from ComCo's refill device to an OEM's refill         device, the ComCo's refill device and the OEM's refill device         must share a common key or a variant key. This key is         fill/refill key with respect to the OEM's refill device and it         is the transfer key with respect to the ComCo's refill device.

For a ComCo to successfully fill/refill its refill device from another refill device (which is higher in the heirachy possibly belonging to the QACo), the ComCo's refill device and the QACo's refill device must share a common key or a variant key. This key is fill/refill key with respect to the ComCo's refill device and it is the transfer key with respect to the QACo's refill device.

26.4.1.1.1 Ink—Remaining Field Information

Table 287 shows the field information for an _(M0) field storing logical ink-remaining amounts in the refill device and which has the ability to transfer down the heirachy.

Attribute Name Value Explanation Type For e.g - Type describing the logical ink TYPE_HIGHQUALITY_BLACK_INK^(a) stored in the ink-remaining field in the refill device. KeyNum Slot number of the refill key. Only the refill key has authenticated ReadWrite access to this field. Non Auth 0 Non authenticated ReadWrite is not RW Perm^(b) allowed to the field. Auth RW 1 Authenticated (key based) ReadWrite Perm^(c) access is allowed to the field. KeyPerm KeyPerms[KeyNum] = 0 KeyNum is the slot number of the refill key, which has ReadWrite permission to the field. KeyPerms[Slot Num of transfer Transfer key can decrement the field. key] = 1 KeyPerms[others = 0 . . . 7 All other keys have ReadOnly access. (except transfer key)] = 0 End Pos Set as required. Depends on the amount of logical ink the device can store and storage resolution - i.e in picolitres or in microlitres. ^(a)This is a sample type only and is not included in the Type Map in Appendix A. ^(b)Non authenticated Read Write permission. ^(c)Authenticated Read Write permission. 26.4.2 Peer to Peer Transfer

Referring to FIG. 381, this transfer is typically performed when ink is transferred from OEM's Ink Refill Device to another Ink Refill Device belonging to the same OEM, or OEM's Ink Refill Device to Ink Device belonging to the same OEM.

26.4.2.1 Keys and Access Permission

There is an ink-remaining field associated with the refill device which transfers ink amounts to other refill devices (peer devices), or to other ink devices. This ink-remaining field has a single key associated with:

-   -   This key transfers ink to the device from another refill device         (which is higher or at the same level in the heirachy),         fills/refills (upgrades) the device itself, and additionally         transfers ink from it to other devices (which are lower in the         heirachy), fills/refills (upgrades) other devices from it.

This key is referred to as the fill/refill key and is used for both fill/refill and transfer. Hence, this key has both ReadWrite and Decrement-Only permission to the ink-remaining field in the refill device.

26.4.2.1.1 Ink-Remaining Field Information

Table 288 shows the field information for an _(M0) field storing logical ink-remaining amounts in the refill device with the ability to transfer between peers.

Attribute Name Value Explanation Type For e.g - Type describing the logical ink TYPE_HIGHQUALITY_BLACK_INK^(a) stored in the ink-remaining field in the refill device. KeyNum Slot number of the refill key. Only the refill key has authen- ticated ReadWrite access to this field. Non Auth 0 Non authenticated ReadWrite is not RW Perm^(b) allowed to the field. Auth RW 1 Authenticated (key based) Perm^(c) ReadWrite access is allowed to the field. KeyPerm KeyPerms[KeyNum] = 1 KeyNum is the slot number of the refill key, which has ReadWrite and Decrement permission to the field. KeyPerms[others = 0 . . . 7 All other keys have ReadOnly (except KeyNum)] = 0 access. End Pos Set as required. Depends on the amount of logical ink the device can store and storage resolution - i.e in picolitres or in microlitres. ^(a)This is a sample type only and is not included in the Type Map in Appendix A. ^(b)Non authenticated Read Write permission. ^(c)Authenticated Read Write permission. 27 Functions 27.1 XferAmount

-   -   Input: KeyRef, _(M0)OfExternal, _(M1)OfExternal, ChipID,         FieldNumL, FieldNumE, XferValLength, XferVal,         InputParameterCheck (optional), R_(E), SIG_(E), R_(E2)     -   Output: ResultFlag, FieldSelect, FieldVal, R_(L2), SIG_(out)     -   Changes: _(M0) and R_(L)     -   Availability Ink refill QA Device         27.1.1 Function Description

The XferAmount function produces data and signature for updating a given _(M0) field. This data and signature when applied to the appropriate device through the WriteFieldsAuth function, will update the _(M0) field of the device.

The system calls the XferAmount function on the upgrade device with a certain XferVal, this XferVal is validated by the XferAmount function for various rules as described in Section 27.1.4, the function then produces the data and signature for the passing into the WriteFieldsAuth function for the device being upgraded.

The transfer amount output consists of the new data for the field being upgraded, field data of the two sequence fields, and a signature using the refill key. When a transfer output is produced, the sequence field data in SEQ_(—)1 is decremented by 2 from the previous value (as passed in with the input), and the sequence field data in SEQ 2 is decremented by 1 from the previous value (as passed in with the input).

Additional InputParameterCheck value must be provided for the parameters not included in the SIG_(E), if the transmission between the System and Ink Refill QA Device is error prone, and these errors are not corrected by the transimission protocol itself. InputParameterCheck is SHA-1[FieldNumL|FieldNumE|XferValLength|XferVal], and is required to ensure the integrity of these parameters, when these inputs are received by the Ink Refill QA Device. This will prevent an incorrect transfer amount being deducted.

The XferAmount function must first calculate the SHA-1[FieldNumL|FieldNumE|XferValLength|XferVal], compare the calculated value to the value received (InputParameterCheck) and only if the values match act upon the inputs.

27.1.2 Input Parameters

Table 289 describes each of the input parameters for XferAmount function.

Parameter Description KeyRef For comsmon key input and output signature: KeyRef.keyNum = Slot number of the key to be used for testing input signature and producing the output signature. SIG_(E) produced using K_(KeyRef.keyNum) by the QA Device being upgraded. SIGout produced using K_(KeyRef.keyNum) for delivery to the QA Device being upgraded. KeyRef.useChipId = 0 For variant key input and output signatures: KeyRef.keyNum = Slot number of the key to be used for generating the variant key. SIG_(E) produced using a variant of K_(KeyRef.keyNum) by the QA Device being upgraded. SIGout produced using a variant of K_(KeyRef.keyNum) for delivery to the QA Device being upgraded. KeyRef.useChipId = 1 KeyRef.chipId = ChipId of the device which generated SIG_(E) and will receive SIGout. _(M0)OfExternal All 16 words of _(M0) of the QA Device being upgraded. _(M1)OfExternal All 16 words of _(M1) of the QA Device being upgraded. ChipId ChipId of the QA Device being upgraded. FieldNumL _(M0) field number of the local (refill) device from which the value will be transferred. FieldNumE _(M0) field number of the QA Device being upgraded to which the value will be transferred. XferValLength XferVal length in words. Non zero length required. XferVal The logical amount that will be transferred from the local device to the external device. R_(E) External random value used to verify input signature. This will be the R from the input signature generator (i.e device generating SIG_(E)). The input signal generator in this case, is the device being upgraded or a translation device. R_(E2) External random value used to produce output signature. This will be R obtained by calling the Random function on the device which will receive the SIG_(out) from the XferAmount function. The device receiving the SIG_(out) in this case, is the device being upgraded or a translation device. SIG_(E) External signature required for authenticating input data. The input data in this case, is the output from the Read function performed on the device being upgraded. A correct SIG_(E) = SIG_(KeyRef)(Data | R_(E) | R_(L)). 27.1.2.1 Input Signature Verification Data Format

The input signature passed in to the XferAmount function is the output signature from the Read function of the Ink QA Device.

FIG. 382 shows the input signature verification data format for the XferAmount function.

Table 290 gives the parameters included in SIG_(E) for XferAmount.

Length Value set Value set Parameter in bits internally from Input RWSense 3 000 Refer to Section 15.3.1.1 MSelect 4 0011 KeyIdSelect 8 00000000 ChipId 48 ChipId of the QA Device being upgraded WordSelect for M₀ 16 All bits set to 1 WordSelect for M₁ 16 All bits set to 1 M0 512 ● M1 512 ● R_(E) 160 ● R_(L) 160 Based on the ● internal R

The XferAmount function is not passed all the parameters required to generate SIG_(E). For producing SIG_(L) which is used to test SIG_(E), the function uses the expected values of some the parameters.

27.1.3 Output Parameters

Table 291 describes each of the output parameters for XferAmount.

Parameter Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Table 47. FieldSelect Selection of fields to be written In this case the bit corresponding to SEQ_1 , SEQ_2 and to FieldNumE are set to 1. All other bits are set to 0. FieldVal Updated data words for Sequence data field and FieldNumE for QA Device being upgraded. Starts with LSW of lower field. This must be passed as input to the WriteFieldsAuth function of the QA Device being upgraded. R_(L2) Internal random value required to generate output signature. This must be passed as input to the WriteFieldsAuth function or Translate function of the QA Device being upgraded. SIG_(out) Output signature which must be passed as an input to the WriteFieldsAuth function of the QA Device being upgraded. SIG_(out) = SIG_(KeyRef)(data | R_(L2) | R_(E2)) as per FIG. 373.

TABLE 292 Result Flag definitions for XferAmount ResultFlag Definition Description FieldNumEInvalid FieldNum to which the amount is being transferred, or which is being upgraded in the QA Device being upgraded is invalid. SeqFieldInvalid The sequence field for the QA Device being upgraded is invalid. FieldNumEWritePermInvalid FieldNum to which the amount is being transferred, or which is being upgraded in the QA Device being upgraded has no authenti- cated write permission. FieldNumLInvalid FieldNum from which the amount is being transferred, or from which the value is being copied in the Upgrading QA Device is invalid. FieldNumLWritePermInvalid FieldNum from which the amount is being transferred in the Upgrading QA Device has no au- thenticated permission, or no authenticated permission with the KeyRef. TypeMismatch Type of the data from which the amount is being transferred in the Upgrading QA Device, doesn't match the Type of data to which the amount in being transferred in the Device being upgraded. UpgradeFieldEInvalid Only applicable for transferring count-remaining values. The upgrade field associated with the count- remaining field in the QA Device being upgraded is invalid. UpgradeFieldLInvalid Only applicable for transferring count-remaining values. The upgrade field associated with the count- remaining field in the Upgrading QA Device is invalid. UpgradeFieldMismatch Only applicable for transferring count-remaining values. Type of the data in the upgrade field in the Upgrading QA Device, doesn't match the Type of data in the upgrade field in the Device being upgraded. FieldNumESizeInsufficient FieldNum to which the amount is being transferred, or which is being upgraded in the QA Device is not big enough to store the trans- ferred data. FieldNumLAmountInsufficient FieldNum in the Upgrading QA Device from which the amount is being transferred doesn't have the amount required for the transfer. 27.1.3.1 SIG_(out)

Refer to Section 20.2.1 for details.

27.1.4 Function Sequence

The XferAmount command is illustrated by the following pseudocode:

Accept input parameters-KeyRef, M0OfExternal, M1OfExternal, ChipId, FieldNumL, FieldNumE, XferValLength # Accept XferVal words For i

0 to XferValLength Accept next XferVal EndFor Accept R_(E), SIG_(E), R_(E2) #Generate message for passing into ValidateKeyRefAndSignature function data

(RWSense|MSelect|KeyIdSelect|ChipId| WordSelect|M0|M1) # Refer to Figure 382. ---------------------------------------------------------------- # Validate KeyRef, and then verify signature ResultFlag = ValidateKeyRefAndSignature(KeyRef,data,R_(E),R_(L)) If (ResultFlag ≠ Pass) Output ResultFlag Return EndIf ---------------------------------------------------------------- #Validate FieldNumE # FieldNumE is present in the device being upgraded PresentFlagFieldNumE

GetFieldPresent(M1OfExternal, FieldNumE) # Check FieldNumE present flag If(PresentFlagFieldNumE ≠ 1) ResultFlag

FieldNumEInvalid Output ResultFlag Return EndIf ------------------------------------------------------------------ --------------- # Check Seq Fields Exist and get their Field Num # Get Seqdata field SEQ_1 num for the device being upgraded XferSEQ_1FieldNum

GetFieldNum(M1OfExternal, SEQ_1) # Check if the Seqdata field SEQ_1 is valid If(XferSEQ_1FieldNum invalid) ResultFlag

SeqFieldInvalid Output ResultFlag Return EndIf # Get Seqdata field SEQ_2 num for the device being upgraded XferSEQ_2FieldNum

GetFieldNum(M1OfExternal, SEQ_2) # Check if the Seqdata field SEQ_2 is valid If(XferSEQ_2FieldNum invalid) ResultFlag

SeqFieldInvalid Output ResultFlag Return EndIf ---------------------------------------------------------------- #Check write permission for FieldNumE PermOKFieldNumE

CheckFieldNumEPerm(M1OfExternal, FieldNumE) If(PermOKFieldNumE ≠ 1) ResultFlag

FieldNumEWritePermInvalid Output ResultFlag Return EndIf ---------------------------------------------------------------- #Check that both SeqData fields have Decrement-Only permission with the same key #that has write permission on FieldNumE PermOKXferSeqData

CheckSeqDataFieldPerms(M1OfExternal, XferSEQ_1FieldNum, XferSEQ_2FieldNum,FieldNumE) If(PermOKXferSeqData ≠ 1) ResultFlag

SeqWritePermInvalid Output ResultFlag Return EndIf ---------------------------------------------------------------- # Get SeqData SEQ_1 data from device being upgraded GetFieldDataWords(XferSEQ_1FieldNum, XferSEQ_1DataFromDevice,M0OfExternal,M1OfExternal) # Get SeqData SEQ_2 data from device being upgraded GetFieldDataWords(XferSEQ_2FieldNum, XferSEQ_2DataFromDevice, M0OfExternal,M1OfExternal) ---------------------------------------------------------------- # FieldNumL is a present in the refill device PresentFlagFieldNumL

GetFieldPresent(M1,FieldNumL) If(PresentFlagFieldNumL ≠ 1) ResultFlag

FieldNumLInvalid Output ResultFlag Return EndIf #Check permission for FieldNumL PermOKFieldNumL

CheckFieldNumLPerm(M1, FieldNumL,KeyRef) If(PermOKFieldNumL ≠ 1) ResultFlag

FieldNumLWritePermInvalid Output ResultFlag Return EndIf ---------------------------------------------------------------- #Find the type attribute for FieldNumE TypeFieldNumE←FindNumType(M1 OfExternal,FieldNumE)

TypeFieldNumE

FindFieldNumType(M1OfExternal,FieldNumE) #Find the type attribute for FieldNumL TypeFieldNumL

FindFieldNumType(M1,FieldNumL) # Check type attribute for both fields match If(TypeFieldNumE ≠TypeFieldNumL) ResultFlag

TypeMismatch Output ResultFlag Return EndIf ---------------------------------------------------------------- ----------------------------------------------------------------------- Do this if the Refill Device is tranferring Count-remaining for Printer upgrades # If the Type is count remaining, check that upgrade values associated with # the count remaining are valid. Refer to Section 28. for further details on # count remaining and upgrade value. If(TypeFieldNumL = TYPE_COUNT_REMAINING)

(TypeFieldNumE =TYPE_COUNT_REMAINING) #Upgrade value field is lower adjoining field UpgradeValueFieldNumE = FieldNumE −1 If(UpgradeValueFieldNumE < 0) # upgrade field doesn't exist for QA Device being upgraded ResultFlag

UpgradeFieldEInvalid Output ResultFlag Return EndIf UpgradeValueFieldNumL = FieldNumL − 1 If(UpgradeValueFieldNumL < 0) # upgrade field doesn't exist for local device ResultFlag

UpgradeFieldLInvalid Output ResultFlag Return EndIf UpgradeValueCheckOK

UpgradeValCheck(UpgradeValueFieldNumL,M0,M1, UpgradeValueFieldNumL,M0OfExternal,M1OfExternal,KeyRef) If(UpgradeValueCheckOK = 0) ResultFlag

UpgradeFieldMismatch Output ResultFlag Return EndIf EndIf # Do this if Field Type is Count Remaining........end ------------------------------------------------------------ #Check whether the device being upgraded can hold the transfer amount #(XferVal + AmountLeft OverFlow

CanHold(FieldNumE,M0OfExternal,XferVal) If OverFlow error ResultFlag

FieldNumESizeInsufficient Output ResultFlag Return EndIf ------------------------------------------------------------- #Check the refill device has the desired amount (XferVal < = AmountLeft) UnderFlow

HasAmount(FieldNumL,M0,XferVal) If UnderFlow error ResultFlag

FieldNumLAmountInsufficient Output ResultFlag Return EndIf ------------------------------------------------------------ # All checks complete ..... # Generate Seqdata for SEQ_1 and SEQ_2 fields XferSEQ_1DataToDevice = XferSEQ_1DataFromDevice − 2 XferSEQ_2DataToDevice = XferSEQ_2DataFromDevice − 1 # Add DataSet to Xfer Entry Cache AddDataSetToXferEntryCache(ChipId,FieldNumE, FieldNumL, XferLength, XferVal, XferSEQ_1DataFromDevice, XferSEQ_2DataFromDevice) # Get current FieldDataE field data words to write to Xfer Entry cache GetFieldDataWords(FieldNumE,FieldDataE,M0OfExternal, M1OfExternal) #Deduct XferVal from FieldNumL and Write new value DeductAndWriteValToFieldNumL(XferVal,FieldNumL,M0) #Generate new field data words for FieldNumE. The current FieldDataE is added to # XferVal to generate new FieldDataE GenerateNewFieldData(FieldNumE,XferVal,FieldDataE) # Generate FieldSelect and FieldVal for SeqData field SEQ_1, SEQ_2 and # FieldDataE... CurrentFieldSelect

0 FieldVal

0 GenerateFieldSelectAndFieldVal(FieldNumE, FieldDataE, XferSEQ_1FieldNum, XferSEQ_1DataToDevice, XferSEQ_2FieldNum, XferSEQ_2DataToDevice, FieldSelect,FieldVal) #Generate message for passing into GenerateSignature function data

(RWSense|FieldSelect|ChipId|FieldVal)# Refer to Figure 373. #Create output signature for FieldNumE SIG_(out)

GenerateSignature(KeyRef,data,R_(L2),R_(E2)) Update R_(L2) to R_(L3) ResultFlag

Pass Output ResultFlag, FieldData, R_(L2 ,) SIG_(out) Return EndIf 27.1.4.1 ResultFlag ValidateKeyRefAndSignature(KeyRef,data,R_(E),R_(L))

This function checks KeyRef is valID, and if KeyRef is valID, then input signature is verified using KeyRef.

CheckRange(KeyRef.keyNum) If invalid ResultFlag

InValidKey Output ResultFlag Return EndIf #Generate message for passing into GenerateSignature function data

(RWSense|MSelect|KeyIdSelect|ChipId|WordSelect|M0|M1) # Refer to Figure 382. #Generate Signature SIG_(L)

GenerateSignature(KeyRef,data,R_(E),R_(L)) # Check input signature SIG_(E) If(SIG_(L) = SIG_(E)) Update R_(L) to R_(L2) Else ResultFlag

Bad Signature Output ResultFlag Return EndIf 27.1.4.2 GenerateFieldSelectAndFieldVal (FieldNumE, FieldDataE,

-   -   XferSEQ_(—)1FieldNum, XferSEQ_(—)1DataToDevice,         XferSEQ_(—)2FieldNum, XferSEQ_(—)2DataToDevice, FieldSelect,         FieldVal)

This functions generates the FieldSelect and FieldVal for output from FieldNumE and its final data, and data to be written to Seq fields SEQ_(—)1 and SEQ_(—)2.

27.1.4.3 PresentFlag GetFieldPresent(M1,FieldNum)

This function checks whether FieldNum is a valid.

FieldSize[16]

0 # Array to hold FieldSize assuming there are 16 fields NumFields

FindNumberOfFieldsInM0(M1,FieldSize) #Refer to Section 19.4.1 If(FieldNum< NumFields) PresentFlag

1 Else PresentFlag

0 EndIf Return PresentFlag 27.1.4.4 NumFields FindNumOfFieldsInM0(M1,FieldSize

)

Refer to Figure 19.4.1 for details.

27.1.4.5 FieldNum GetFieldNum(M1, Type)

This function returns the field number based on the Type.

FieldSize[16]

0 # Array to hold FieldSize assuming there are 16 fields NumFields

FindNumberOfFieldsInM0(M1,FieldSize) #Refer to Section 19.4.1 For i = 0 to NumFields If(M1[i].Type = Type) Return i # This is field Num for matching field EndFor i = 255 # If XferSession field was not found then return an invalid value Return i 27.1.4.6 PermOK CheckFieldNumEPerm(M1,FieldNumE)

This function checks authenticated write permission for FieldNum which holds the upgraded value.

AuthRW

M1[FieldNum].AuthRW NonAuthRW

M1[FieldNum].NonAuthRW If(AuthRW = 1)

(NonAuthRW = 0) PermOK

1 Else PermOK

0 EndIf Return PermOK 27.1.4.7 PermOK CheckSeqDataFieldPerms(M1, XferSEQ_(—)1FieldNum, XferSEQ_(—)2FieldNum, FieldNumE)

This function checks that both SeqData fields have Decrement-Only permission with the same key that has write permission on FieldNumE.

KeyNumForFieldNumE

M1[FieldNumE].KeyNum # Isolate KeyNum for the field that will # be upgraded # Isolate KeyNum for both SeqData fields and check that they can be written using the same key KeyNumForSEQ_1

M1[XferSEQ_1FieldNum].KeyNum KeyNumForSEQ_2

M1[XferSEQ_2FieldNum].KeyNum If(KeyNumForSEQ_1 ≠KeyNumForSEQ_2) PermOK

0 Return PermOK EndIf # Check that the write key for FieldNumE and SeqData field is not the same If (KeyNumForSEQ_1 = KeyNumForFieldNumE) PermOK

0 Return PermOK EndIf #Isolate Decrement-Only permissions with the write key of FieldNumE KeyPermsSEQ_1

M1[XferSEQ_1FieldNum].KeyPerms[KeyNumForFieldNumE] KeyPermsSEQ_2

M1[XferSEQ_2FieldNum].KeyPerms[KeyNumForFieldNumE] # Check that both sequence fields have Decrement-Only permission for this key If(KeyPermsSEQ_1 =0)

(KeyPermsSEQ_2 = 0) PermOK

0 Return PermOK EndIf PermOK

1 Return PermOK 27.1.4.8 AddDataSetToXferEntryCache (ChipID, FieldNumE, FieldNumL, XferVal, SEQ_(—)1Data, SEQ_(—)2Data)

This function adds a new dataset to the Xfer Entry cache. Dataset is a single record in the Xfer Entrycache. Refer to Section 27 for details.

# Search for matching ChipId FieldNumE is Cache DataSet

SearchDataSetInCache (ChipId, FieldNumE) # If found If(DataSet is valid) DeleteDataSetInCache(DataSet) # This creates a vacant dataset AddRecordToCache(ChipId, FieldNumE,FieldDataL, XferVal,SEQ_1Data, SEQ_2Data) EndIf # Searches the cache for XferState complete/deleted Found

SearchRecordsInCache(complete/deleted) If(Found =1) AddRecordToCache(ChipId, FieldNumE,FieldDataL, XferVal,SEQ_1Data, SEQ_2Data) Else # This will overwrite the oldest DataSet in cache AddRecordToCache(ChipId, FieldNumE,FieldDataL, XferVal,SEQ_1Data, SEQ_2Data) Return Endif Set XferState in record to Xfer Return 27.1.4.9 FieldType FindFieldNumType(M1,FieldNum)

This function gets the Type attribute for a given field.

-   -   FieldType←M1[FieldNum].Type     -   Return FieldType         27.1.4.10 PermOK CheckFieldNumLPerm(M1, FieldNumL, KeyRef)

This function checks authenticated write permissions using KeyRef for FieldNumL in the refill device.

AuthRW

_(M1)[FieldNumL].AuthRW KeyNumAtt

_(M1)[FieldNumL].KeyNum DOForKeys

_(M1)[FieldNumL].DOForKeys[KeyNum] # Authenticated write allowed # ReadWrite key for field is the same as Input KeyRef.keyNum # Key has both ReadWrite and DecrementOnly Permission If(AuthRW = 1)

(KeyRef.keyNum = KeyNumAtt)

(DOForKeys = 1 PermOK

1 Else PermOK

0 EndIf Return PermOK 27.1.4.11 CheckOK UpgradeValCheck(FieldNum1, M0OfFieldNum1, M1OfFieldNum1, FieldNum2, M0OfFieldNum2, M1OfFieldNum2,KeyRef)

This function checks the upgrade value corresponding to the count remaining. The upgrade value corresponding to the count remaining field is stored in the lower adjoining field. To upgrade the count remaining field, the upgrade value in refill device and the device being upgraded must match.

#Check authenticated write permissions is allowed to the field #Check that only one key has ReadWrite access, #and all other keys are ReadOnly access PermCheckOKFieldNum1

CheckUpgradeKeyForField(FieldNum1,M1OfFieldNum1,KeyRef) If(PermCheckOKFieldNum1 ≠ 1) CheckOK

0 Return CheckOK EndIf PermCheckOKFieldNum2

CheckUpgradeKeyForField(FieldNum2,M1OfFieldNum2,KeyRef) If(PermCheckOKFieldNum2 ≠ 1) CheckOK

0 Return CheckOK EndIf #Get the upgrade value associated with field GetFieldDataWords(FieldNum1,UpgradeValueFieldNum1, M0OfFieldNum1, M1OfFieldNum1) #Get the upgrade value associated with field GetFieldDataWords(FieldNum2,UpgradeValueFieldNum2, M0OfFieldNum2, M1OfFieldNum2) If(UpgradeValueFieldNum1 ≠ UpgradeValueFieldNum2) CheckOK

0 Return CheckOK EndIf # Get the type attribute for the field UpgradeTypeFieldNum1

GetUpgradeType(FieldNum1,M1OfFieldNum1) UpgradeTypeFieldNum2

GetUpgradeType(FieldNum2,M1OfFieldNum2) If(UpgradeTypeFieldNum1 ≠ UpgradeTypeFieldNum2) CheckOK

0 Return CheckOK EndIf CheckOK

1 Return CheckOK 27.1.4.12 CheckOK CheckUpgradeKeyForField(FieldNum,M1,KeyRef)

This function checks that authenticated write permissions is allowed to the field. It also checks that only one key has ReadWrite access and all other keys have ReadOnly access. KeyRef which updates count remaining must not have write access to the upgarde value field.

KeyNum

M1[FieldNum].KeyNum AuthRW

M1[FieldNum].AuthRW NonAuthRW

M1[FieldNum].NonAuthRW DOForKeys

M1[FieldNum].DOForKeys #Check that KeyRef doesn't have write permissions to the field If(KeyRef.keyNum = KeyNum) CheckOK

0 Return CheckOK EndIf #AuthRW access allowed or NonAuthRW not allowed If(AuthRW = 0)

(NonAuthRW =1) CheckOK

0 Return CheckOK EndIf For i

0 to 7 # Keys other than KeyNum are allowed ReadOnly access, # DecrementOnly access not allowed for other keys(not KeyNum) If (i ≠KeyNum)

(DOForKeys[i] = 1) CheckOK

0 Return CheckOK EndIf #ReadWrite access allowed for KeyNum, #ReadWrite and DecrementOnly access not allowed for KeyNum. If (i = KeyNum)

(DOForKeys[i] = 1) CheckOK

0 Return CheckOK EndIf EndFor CheckOK

1 Return CheckOK 27.1.4.13 UpgradeType GetUpgradeType(FieldNum, M1)

-   -   This function gets the type attribute for the upgrade field.

UpgradeType GetUpgradeType(FieldNum) UpgradeType

M1[FieldNum].Type Return UpgradeType 27.1.4.14 GetFieldDataWords(FieldNum, FieldData

, M0, M1)

This function gets the words corresponding to a given field.

CurrPos

MaxWordInM If FieldNum = 0 CurrPos

MaxWordInM Else CurrPos

(M1[FieldNum −1].EndPos) −1 # Next lower word after last word of the # previous field EndIf EndPos

(M1[FieldNum].EndPos) For i

EndPos to CurrPos j

0 FieldData[j]

M0[i] #Copy M0 word to FieldData array EndFor 27.2 StartRollBack

-   -   Input: KeyRef, _(M0)OfExternal, _(M1)OfExternal, ChipID,         FieldNumL, FieldNumE, InputParameterCheck (optional), R_(E),         SIG_(E), R_(E2)     -   Output ResultFlag, FieldSelect, FieldVal, R_(L2) SIG_(out)     -   Changes: _(M0) and R_(L)     -   Availability Ink refill QA Device and Parameter Upgrader QA         Device         27.2.1 Function Description

StartRollBack function is used to start a rollback sequence if the QA Device being upgraded didn't receive the transfer message correctly and hence didn't receive the transfer.

The system calls the function on the upgrading QA Device, passing in FieldNumE and ChipId of the QA Device being upgraded, and FieldNumL of the upgrading QA Device. The upgrading QA Device checks that the QA Device being upgraded didn't actually receive the message correctly, by comparing the values read from the device with the values stored in the Xfer Entry cache. The values compared is the value of the sequence fields. After all checks are fulfilled, the upgrading QA Device produces the new data for the sequence fields and a signature. This is subsequently applied to the QA Device being upgraded (using the WriteFieldAuth function), which updates the sequence fields SEQ_(—)1 and SEQ_(—)2 to the pre-rollback values. However, the new data for the sequence fields and signature can only be applied if the previous data for the sequence fields produced by Xfer function has not been written.

The output from the StartRollBack function consists only of the field data of the two sequence fields, and a signature using the refill key. When a pre-rollback output is produced, then sequence field data in SEQ_(—)1 (as stored in the Xfer Entry cache, which is what is passed in to the XferAmount function) is decremented by 1 and the sequence field data in SEQ_(—)2 (as stored in the Xfer Entry cache, which is what is passed in to the XferAmount function) is decremented by 2.

Additional InputParameterCheck value must be provided for the parameters not included in the SIG_(E), if the transmission between the System and Ink Refill QA Device is error prone, and these errors are not corrected by the transimission protocol itself. InputParameterCheck is SHA-1[FieldNumL|FieldNumE], and is required to ensure the integrity of these parameters, when these inputs are received by the Ink Refill QA Device.

The StartRollBack function must first calculate the SHA-1[FieldNumL|FieldNumE], compare the calculated value to the value received (InputParameterCheck) and only if the values match act upon the inputs.

27.2.2 Input Parameters

Table 293 describes each of the input parameters for StartRollBack function.

Parameter Description KeyRef For common key input signature: KeyRef.keyNum = Slot number of the key to be used for testing input signature. SIG_(E) produced using K_(KeyRef.keyNum) by the QA Device being upgraded. KeyRef.useChipId = 0 For variant key input signature: KeyRef.keyNum = Slot number of the key to be used for generating the variant key for testing input signature. SIG_(E) produced using a variant of K_(KeyRef.keyNum) by the QA Device being upgraded. KeyRef.useChipId = 1 KeyRef.chipId = ChipId of the device which generated SIG_(E). _(M0)OfExternal All 16 words of _(M0) of the QA Device being upgraded which failed to upgrade. _(M1)OfExternal All 16 words of _(M1) of the QA Device being upgraded which failed to upgrade. ChipId ChipId of the QA Device being upgraded which failed to upgrade. FieldNumL _(M0) field number of the local (refill) device from which the value was supposed to transferred. FieldNumE _(M0) field number of the QA Device being upgraded to which the value couldn't be transferred. R_(E) External random value used to verify input signature. This will be the R from the input signature genera- tor (i.e device generating SIG_(E)). The input signal generator in this case, is the device which failed to upgrade or a translation device. SIG_(E) External signature required for authenticating input data. The input data in this case, is the output from the Read function performed on the device which failed to upgrade. A correct SIG_(E) = SIG_(KeyRef)(Data | R_(E) | R_(L)). 27.2.2.1 Input Signature Verification Data Format

Refer to Section 27.1.2.1.

27.2.3 Output Parameters

Table 294 describes each of the output parameters for StartRollBack function.

Parameter Description ResultFlag Indicates whether the function completed success- fully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1, Table 292 and Table 295. FieldSelect Selection of fields to be written In this case the bits corresponding to SEQ_1 and SEQ_2 are set to 1. All other bits are set to 0. FieldVal Updated data for sequence datat field for QA Device being upgraded. This must be passed as input to the WriteFieldsAuth function of the QA Device being upgraded. R_(L2) Internal random value required to generate output signature. This must be passed as input to the WriteFieldsAuth function or Translate function of the QA Device being upgraded. SIG_(out) Output signature which must be passed as an input to the WriteFieldsAuth function of the QA Device being upgraded. SIG_(out) = SIG_(KeyRef)(data | R_(L2) | R_(E2)) as per FIG. 373.

TABLE 295 Result definition for StartRollBack ResultFlag Definition Description RollBackInvalid RollBack cannot be performed on the request because parameters for rollback is incorrect. 27.2.3.1 SIG_(out)

Refer to Section 20.2.1 for details.

27.2.4 Function Sequence

The StartRollBack command is illustrated by the following pseudocode:

-   -   Accept input parameters—KeyRef, M1OfExternal, ChipId, FieldNumL,         FieldNumE, R_(E), SIG_(E), R_(E2)

Accept R_(E), SIG_(E), R_(E2) #Generate message for passing into ValidateKeyRefAndSignature function data

(RWSense|MSelect|KeyIdSelect|ChipId|WordSelect|M0|M1) # Refer to Figure 382. ---------------------------------------------------------------- # Validate KeyRef, and then verify signature ResultFlag = ValidateKeyRefAndSignature(KeyRef,data,R_(E),R_(L)) If (ResultFlag ≠ Pass) Output ResultFlag Return EndIf ----------------------------------------------------------------# Check Seq Fields Exist and get their Field Num # Get Seqdata field SEQ_1 num for the device being upgraded XferSEQ_1FieldNum

GetFieldNum(M1OfExternal, SEQ_1) # Check if the Seqdata field SEQ_1 is valid If(XferSEQ_1FieldNum invalid) ResultFlag

SeqFieldInvalid Output ResultFlag Return EndIf # Get Seqdata field SEQ_2 num for the device being upgraded XferSEQ_2FieldNum

GetFieldNum(M1OfExternal, SEQ_2) # Check if the Seqdata field SEQ_2 is valid If(XferSEQ_2FieldNum invalid) ResultFlag

SeqFieldInvalid Output ResultFlag Return EndIf ---------------------------------------------------------------- # Get SeqData SEQ_1 data from device being upgraded GetFieldDataWords(XferSEQ_1FieldNum, XferSEQ_1DataFromDevice,M0OfExternal,M1OfExternal) # Get SeqData SEQ_2 data from device being upgraded GetFieldDataWords(XferSEQ_2FieldNum, XferSEQ_2DataFromDevice, M0OfExternal,M1OfExternal) ---------------------------------------------------------------- # Check Xfer Entry in cache is correct - dataset exists, Field data # and sequence field data matches and Xfer State is correct XferEntryOK

CheckEntry(ChipId, FieldNumE, FieldNumL, XferSEQ_1DataFromDevice, XferSEQ_2DataFromDevice) If( XferEntryOK= 0) ResultFlag

RollBackInvalid Output ResultFlag Return EndIf # Generate Seqdata for SEQ_1 and SEQ_2 fields XferSEQ_1DataToDevice = XferSEQ_1DataFromDevice − 1 XferSEQ_2DataToDevice = XferSEQ_2DataFromDevice − 2 # Generate FieldSelect and FieldVal for sequence fields SEQ_1 and SEQ_2 CurrentFieldSelect

0 FieldVal

0 GenerateFieldSelectAndFieldVal(XferSEQ_1FieldNum, XferSEQ_1DataToDevice, XferSEQ_2FieldNum, XferSEQ_2DataToDevice, FieldSelect, FieldVal) #Generate message for passing into GenerateSignature function data

(RWSense|FieldSelect|ChipId|FieldVal)# Refer to Figure 373. #Create output signature for FieldNumE SIG_(out)

GenerateSignature(KeyRef,data,R_(L2),R_(E2)) Update R_(L2) to R_(L3) ResultFlag

Pass Output ResultFlag, FieldData, R_(L2) ,SIG_(out) Return EndIf 27.3 RollBackAmount

-   -   Input KeyRef, _(M0)OfExternal, _(M1)OfExternal, ChipID,         FieldNumL, FieldNumE, InputParameterCheck (optional), R_(E),         SIG_(E)     -   Output: ResultFlag     -   Changes: _(M0) and R_(L)     -   Availablity: Ink refill QA Device         27.3.1 Function Description

RollBackAmount function finally adjusts the value of the FieldNumL of the upgarding QA Device to a previous value before the transfer request, if the QA Device being upgraded didn't receive the transfer message correctly (and hence was not upgraded).

The upgrading QA Device checks that the QA Device being upgraded didn't actually receive the transfer message correctly, by comparing the sequence data field values read from the device with the values stored in the Xfer Entry cache. The sequence data field values read must match what was previously written using the StartRollBack function. After all checks are fulfilled, the upgrading QA Device adjusts its FieldNumL.

Additional InputParameterCheck value must be provided for the parameters not included in the SIG_(E), if the transmission between the System and Ink Refill QA Device is error prone, and these errors are not corrected by the transimission protocol itself. InputParameterCheck is SHA-1[FieldNumL|FieldNumE], and is required to ensure the integrity of these parameters, when these inputs are received by the Ink Refill QA Device.

The RollBackAmount function must first calculate the SHA-1[FieldNumL|FieldNumE], compare the calculated value to the value received (InputParameterCheck) and only if the values match act upon the inputs.

27.3.2 Input Parameters

Table 296 describes each of the intput parameters for RollbackAmount function.

Parameter Description KeyRef For common key input signature: KeyRef.keyNum = Slot number of the key to be used for testing input signature. SIG_(E) produced using K_(KeyRef.keyNum) by the QA Device being upgraded. KeyRef.useChipId = 0 For variant key input signature: KeyRef.keyNum = Slot number of the key to be used for generating the variant key for testing input signature. SIG_(E) produced using a variant of K_(KeyRef.keyNum) by the QA Device being upgraded. KeyRef.useChipId = 1 KeyRef.chipId = ChipId of the device which generated SIG_(E). _(M0)OfExternal All 16 words of _(M0) of the QA Device being upgraded which failed to upgrade. _(M1)OfExternal All 16 words of _(M1) of the QA Device being upgraded which failed to upgrade. ChipId ChipId of the QA Device being upgraded which failed to upgrade. FieldNumL _(M0) field number of the local (refill) device from which the value was supposed to transferred. FieldNumE _(M0) field number of the QA Device being upgraded to which the value was not transferred. R_(E) External random value used to verify input signature. This will be the R from the input signature generator (i.e device generating SIG_(E)). The input signal generator in this case, is the device which failed to upgrade or a translation device. SIG_(E) External signature required for authenti- cating input data. The input data in this case, is the output from the Read function performed on the device which failed to upgrade. A correct SIG_(E) = SIG_(KeyRef)(Data | R_(E) | R_(L)). 27.3.2.1 Input Signature Generation Data Format

Refer to Section 27.1.2.1 for details.

27.3.3 Output Parameters

Table 297 describes each of the output parameters for RollbackAmount.

Parameter Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1, Table 292 and Table 295. 27.3.4 Function Sequence

The RollBackAmount command is illustrated by the following pseudocode:

Accept input parameters-KeyRef, M0OfExternal, M1OfExternal, ChipId, FieldNumL, FieldNumE, R_(E),SIG_(E) #Generate message for passing into ValidateKeyRefAndSignature function data

(RWSense|MSelect|KeyIdSelect|ChipId|WordSelect|M0|M1) # Refer to Figure 382. ---------------------------------------------------------------- # Validate KeyRef, and then verify signature ResultFlag = ValidateKeyRefAndSignature(KeyRef,data,R_(E),R_(L)) If (ResultFlag ≠ Pass) Output ResultFlag Return EndIf ---------------------------------------------------------------- # Check Seq Fields Exist and get their Field Num # Get Seqdata field SEQ_1 num for the device being upgraded XferSEQ_1FieldNum

GetFieldNum(M1OfExternal, SEQ_1) # Check if the Seqdata field SEQ_1 is valid If(XferSEQ_1FieldNum invalid) ResultFlag

SeqFieldInvalid Output ResultFlag Return EndIf # Get Seqdata field SEQ_2 num for the device being upgraded XferSEQ_2FieldNum

GetFieldNum(M1OfExternal, SEQ_2) # Check if the Seqdata field SEQ_2 is valid If(XferSEQ_2FieldNum invalid) ResultFlag

SeqFieldInvalid Output ResultFlag Return EndIf ---------------------------------------------------------------- # Get SeqData SEQ_1 data from device being upgraded GetFieldDataWords(XferSEQ_1FieldNum, XferSEQ_1DataFromDevice,M0OfExternal,M1OfExternal) # Get SeqData SEQ_2 data from device being upgraded GetFieldDataWords(XferSEQ_2FieldNum, XferSEQ_2DataFromDevice, M0OfExternal,M1OfExternal) ---------------------------------------------------------------- # Generate Seqdata for SEQ_1 and SEQ_2 fields with the data that is read XferSEQ_1Data = XferSEQ_1DataFromDevice + 1 XferSEQ_2Data = XferSEQ_2DataFromDevice + 2 # Check Xfer Entry in cache is correct - dataset exists, Field data # and sequence field data matches and Xfer State is correct XferEntryOK

CheckEntry(ChipId, FieldNumE, FieldNumL, XferSEQ_1Data, XferSEQ_2Data) If( XferEntryOK= 0) ResultFlag

RollBackInvalid Output ResultFlag Return EndIf # Get ΔFieldDataL from DataSet GetVal(ChipId, FieldNumE,ΔFieldDataL) # Add ΔFieldDataL to FieldNumL AddValToField(FieldNumL,ΔFieldDataL) # Update XferState in DataSet to complete/deleted UpdateXferStateToComplete(ChipId, FieldNumE) ResultFlag

Pass Output ResultFlag Return Functions Upgrade Device (Printer Upgrade) 28 Concepts

This section is very similar to Section 26. The differences between this section and Section 26 have been summarised and underlined, where required.

28.1 Purpose

In a printing application, a printer contains a Printer QA Device, which stores details of the various operating parameters of a printer, some of which may be upgradeable. The upgradeable parameters must be written (initially) and changed in an authorised manner.

The authorisation for the write or change is achieved by using a Parameter Upgrader QA Device which contains the necessary functions to allow a write or a change of a parameter value (e.g. a print speed) into another QA Device, typically a printer QA Device. This QA Device is also referred to as an upgrading QA Device.

A parameter upgrader QA Device is able to perform a fixed number of upgrades, and this number is effectively a consumable value. The number of upgrades remaining is also referred to as count-remaining. With each write/change of an operating parameter in a Printer QA Device, the count-remaining decreases by 1, and can be replenished by a value upgrader QA Device.

The Parameter Upgrader QA Device can also be referred to as the Upgrading QA Device, and the Printer QA Device can also be referred to as the QA Device being upgraded.

The writing or changing of the parameter can also be referred to as a transfer of a parameter.

The Parameter Upgrader QA Device copies its parameter value field to the parameter value field of Printer QA Device, and decrements the count-remaining field associated with the parameter value field by 1.

28.2 Requirements

The transfer of a parameter has two basic requirements:

-   -   The transfer can only be performed if the transfer request is         valid. The validity of the transfer request must be completely         checked by the Parameter Upgrader QA Device, before it produces         the required output for the transfer. It must not be possible to         apply the transfer output to the Printer QA Device, if the         Parameter Upgrader QA Device has been already been rolled back         for that particular transfer.     -   A process of rollback is available if the transfer was not         received by the Printer QA Device.

A rollback is performed only if the rollback request is valid. The validity of the rollback request must be completely checked by the Parameter Upgrader QA Device, before the count-remaining value is incremented by 1. It must not be possible to rollback an Parameter Upgrader QA Device for a transfer, which has already been applied to the Printer QA Device i.e the Parameter Upgrader QA Device must only be rolled back for transfers that have actually failed.

28.3 Basic Scheme

The transfer and rollback process is shown in FIG. 383.

Following is a sequential description of the transfer and rollback process:

-   1. The System Reads the memory vectors M0 and M1 of the Printer QA     Device. The output from the read which includes the M0 and M1 words     of the Printer QA Device, and a signature, is passed as an input to     the Transfer Request. It is essential that M0 and M1 are read     together. This ensures that the field information for M0 fields are     correct, and have not been modified, or substituted from another     device. Entire M0 and M1 must be read to verify the correctness of     the subsequent Transfer Request by the Parameter Upgrader QA Device. -   2. The System makes a Transfer Request to the Parameter Upgrader QA     Device with the field in the Parameter Upgrader QA Device whose data     will be copied to the Printer QA Device, and the field in Printer QA     Device to which this data will be copied to. The Transfer Request     also includes the output from Read of the Printer QA Device. The     Parameter Upgrader QA Device validates the Transfer Request based on     the Read output, checks that it has enough count-remaining for a     successful transfer, and then produces the necessary Transfer     output. The Transfer Output typically consists of new field data for     the field being refilled or upgraded, additional field data required     to ensure the correctness of transfer/rollback, along with a     signature. -   3. The System then applies the Transfer Output on the Printer QA     Device, by calling an authenticated Write on it, passing in the     Transfer Output. The Write is either successful or not. If the Write     is not successful, then the System will repeat calling the Write     function using the same transfer output, which may be successful or     not. If unsuccessful the System will initiate a rollback of the     transfer. The rollback must be performed on the Parameter Upgrader     QA Device, so that it can adjust its value to a previous value     before the current Transfer Request was initiated. -   4. The System starts a rollback by Reading the memory vectors M0 and     M1 of the Printer QA Device. -   5. The System makes a StartRollBack Request to the Parameter     Upgrader QA Device with same input parameters as the Transfer     Request, and the output from Read in (4). The Parameter Upgrader QA     Device validates the StartRollBack Request based on the Read output,     and then produces the necessary Pre-rollback output. The     Pre-rollback output typically consists only of additional field data     along with a signature. -   6. The System then applies the Pre-rollback output on the Parameter     Upgrader QA Device, by calling an authenticated Write on it, passing     in the Pre-rollback output. The Write is either successful or not.     If the Write is not successful, then either (6), or (5) and (6) must     be repeated. -   7. The System then Reads the memory vectors M0 and M1 of the Printer     QA Device. -   8. The System makes a RollBack Request to the Parameter Upgrader QA     Device with same input parameters as the Transfer Request, and the     output from Read (7). The Parameter Upgrader QA Device validates the     RollBack Request based on the Read output, and then rolls back its     count-remaining field by incrementing it by 1.     28.3.1 Transfer

The Printer QA Device stores upgradeable operating parameter values in M0 fields, and its corresponding M₁ words contains field information for its operating parameter fields. The field information consists of the size of the field, the Type of data stored in field and the access permission to the field. See Section 8.1.1 for details.

The Parameter Upgrader QA Device also stores the new operating parameter values (which will be written to the Printer QA Device) in its M0 fields, and its coressponding M₁ words contains field information for the new operating parameter fields. Additionally, the Parameter Upgrader QA Device has a count-remaining field associated with the new operating parameter value field. The count-remaining field occupies the higher field position when compared to its associated operating parameter value field.

28.3.1.1 Authorisation

The basic authorisation for a transfer comes from a key, which has authenticated ReadWrite permission (stored in field information as KeyNum) to the operating parameter field in the Printer QA Device. We will refer to this key as the upgrade key. The same upgrade key must also have authenticated decrement-only permission to the count-remaining field (which decrements by 1 with every transfer) in the Parameter Upgrader QA Device.

After validating the input upgrade request, the Parameter Upgrader QA Device will decrement the value of the count-remaining field by 1, and produce data (by copying the data stored from its operating parameter field) and signature for the new operating parameter using the upgrade key. Note that the Parameter Upgrader QA Device can decrement its count-remaining field only if the upgrade key has the permission to decrement it.

The data and signature produced by the Parameter Upgrader QA Device is subsequently applied to the Printer QA Device. The Printer QA Device will accept the new transferred operating parameter, only if the signature is valid. Note that the signature will only be valid if it was produced using the upgrade key which has write permission to the operating parameter field being written.

The upgrade key has authenticated ReadWrite permission to the operating parameter field (which will change) in the Printer QA Device. The upgrade key has decrement-only permission to the the count-remaining field (which decrements by 1 with every transfer of field) in the Parameter Upgrader QA Device.

28.3.1.2 Data Type Matching

The Parameter Upgrader QA Device validates the transfer request by matching the Type of the data in the field information of operating parameter field (stored in M1) of Printer QA Device to the Type of data in the field information of operating parameter field of the Parameter Upgrader QA Device. This ensures that equivalent data types are being transferred i.e Network_OEM1_printspeed_(—)1500 is not transferred to Network_OEM1_printspeed_(—)2000.

28.3.1.3 Addition Validation

Additional validation of the transfer request must be performed before a transfer output is generated by the Parameter Upgrader QA Device. These are as follows:

-   -   For the Printer QA Device

-   1. Whether the field being upgraded is actually present.

-   2. Whether the field being upgraded can hold the changed value.     -   For the Parameter Upgrader QA Device:

-   1. Whether the new operating parameter field and its associated     count-remaining is actually present.

-   2. Whether the count-remaining field has an upgrade left for the     transfer to succeed.     28.3.1.4 Rollback Facilitation

To facilitate a rollback, the Parameter Upgrade QA Device will store a list of transfer requests processed by it. This list is referred to as the Xfer Entry cache. Each record in the list consists of the transfer parameters corresponding to the transfer request.

28.3.2 Rollback

A rollback request will be validated by looking through the Xfer Entry cache of the Parameter Upgrader QA Device. After the right transfer request is found the Parameter Upgrade QA Device checks that the output from the transfer request was not applied to the Printer QA Device by comparing the current Read of the Printer QA Device to the values in the Xfer Entry cache, and finally rolling back the Parameter Upgrader QA Device count-remaining field by incrementing it by 1. The Parameter Upgrader QA Device must be absolutely sure that the Printer QA Device didn't receive the transfer. This factor determines the additional fields that must be written along with new operating parameter data, and also the parameters of the transfer request that must be stored in the Xfer Entry cache to facilitate a rollback, to prove that the Printer QA Device didn't actually receive the transfer.

The rollback process increments the count-remaining field by 1 in the Parameter Upgrader QA Device.

28.3.2.1 Sequence Fields

The rollback process must ensure that the transfer output (which was previously produced) for which the rollback is being performed, cannot be applied after the rollback has been performed. How do we achieve this? There are two separate decrement-only sequence fields (SEQ_(—)1 and SEQ_(—)2) in the Printer QA Device which can only be decremented by the Parameter Upgrader QA Device using the upgrade key. The nature of data to be written to the sequence fields is such that either the transfer output or the pre-rollback output can be applied to the Printer QA Device, but not both i.e they must be mutually exclusive. Refer to Table 285 for details.

The two sequence fields are initialised to 1xFFFFFFFF using sequence key. The sequence key is different to the upgrade key, and has authenticated ReadWrite permission to both the sequence fields.

The transfer output consists of the new data for the field being upgraded, field data of the two sequence fields, and a signature using the upgrade key. The field data for SEQ_(—)1 is decremented by 2 from the original value that was passed in with the transfer request. The field data for SEQ_(—)2 is decremented by 1 from the original value that was passed in with the transfer request.

The pre-rollback output consists only of the field data for the two sequence fields, and a signature using the upgrade key. The field data for SEQ_(—)1 is decremented by 1 from the original value that was passed in with the transfer request. The field data for SEQ_(—)2 is decremented by 2 from the original value that was passed in with the transfer request.

Since the two sequence fields are decrement-only fields, the writing of the transfer output to QA Device being upgraded will prevent the writing of the pre-rollback output to QA Device being upgraded, since the sequence fields are decrement-only fields, and only one possible set can be written. If the writing of the transfer output fails, then pre-rollback can be written. However, the transfer output cannot be written after the pre-rollback output has been written.

Before a rollback is performed, the Parameter Upgrader QA Device must confirm that the sequence fields was successfully written to the pre-rollback values in the Printer QA Device. Because the sequence fields are decrement-only fields, the Printer QA Device will allow pre-rollback output to be written only if the transfer output has not been written.

28.3.2.1.1 Field Information of the Sequence Data Field

For a device to be upgradeable the device must have two sequence fields SEQ_(—)1 and SEQ_(—)2 which are written with sequence data during the transfer sequence. Thus all upgrading QA Devices, ink QA Devices and printer QA Devices must have two sequence fields. The upgrading QA Devices must have these fields because they can be upgraded as well. The sequence field information are defined in Table 298.

Attribute Name Value Explanation Type TYPE_SEQ_1 or See Appendix A for exact TYPE_SEQ_2. data. KeyNum Slot number of Only the sequence key the sequence has authenticated ReadWrite key. access to this field. Non Auth 0 Non authenticated ReadWrite RW Perm^(b) is not allowed to the field. Auth RW 1 Authenticated (key based) Perm^(c) ReadWrite access is allowed to the field. KeyPerm KeyPerms[KeyNum] = KeyNum is the slot number 0 of the sequence key, which has ReadWrite permission to the field. KeyPerms[Slot Upgrade key can decrement number of upgrade the sequence field. key] = 1 KeyPerms[others = All other keys have 0 . . . 7(except ReadOnly access. upgrade key)] = 0 End Pos Set as required. Size is typically 1 word. ^(a)This is a sample type only and is not included in the Type Map in Appendix A. ^(b)Non authenticated Read Write permission. ^(c)Authenticated Read Write permission. 28.3.3 Upgrade States

There are three states in an transfer sequence, the first state is initiated for every transfer, while the next two states are initiated only when the transfer fails. The states are—Xfer, StartRollback, and Rollback.

28.3.3.1 Upgrade Flow

FIG. 384 shows a typical upgrade flow.

28.3.3.2 Xfer

This state indicates the start of the transfer process, and is the only state required if the transfer is successful. During this state, the Parameter Upgrader QA Device adds a new record to its Xfer Entry cache, decrements its count-remaining by 1, produces new operating parameter field, new sequence data (as described in Section 28.3.2.1) and a signature based on the upgrade key.

The Printer QA Device will subsequently write the new operating parameter field and new sequence data, after verifying the signature. If the new operating parameter field can be successfully written to the Printer QA Device, then this will finish a successful transfer.

If the writing of the new amount is unsuccessful (result returned is BAD SIG), the System will re-transmit the transfer output to the Printer QA Device, by calling the authenticated Write function on it again, using the same transfer output.

If retrying to write the same transfer output fails repeatedly, the System will start the rollback process on Parameter Upgrader QA Device, by calling the Read function on the Printer QA Device, and subsequently calling the StartRollBack function on the Parameter Upgrader QA Device. After a successful rollback is performed, the System will invoke the transfer sequence again.

28.3.3.3 StartRollBack

This state indicates the start of the rollback process. During this state, the Parameter Upgrade QA Device produces the next sequence data and a signature based on the upgrade key. This is also called a pre-rollback, as described in Section 26.3.2.

The pre-rollback output can only be written to the Printer QA Device, if the previous transfer output has not been written. The writing of the pre-rollback sequence data also ensures, that if the previous transfer output was captured and not applied, then it cannot be applied to the Printer QA Device in the future.

If the writing of the pre-rollback output is unsuccessful (result returned is BAD SIG), the System will re-transmit the pre-rollback output to the Printer QA Device, by calling the authenticated Write function on it again, using the same pre-rollback output.

If retrying to write the same pre-rollback output fails repeatedly, the System will call the StartRollback on the Parameter Upgrade QA Device again, and subsequently calling the authenticated Write function on the Printer QA Device using this output.

28.3.3.4 Rollback

This state indicates a successful deletion (completion) of a transfer sequence. During this state, the Parameter Upgrader QA Device verifies the sequence data produced from StartRollBack has been correctly written to Printer QA Device, then rolls its count-remaining field to a previous value before the transfer request was issued.

28.3.4 Xfer Entry Cache

The Xfer Entry data structure must allow for the following:

-   -   Stores the transfer state and sequence data for a given transfer         sequence.     -   Store all data corresponding to a given transfer, to facilitate         a rollback to the previous value before the transfer output was         generated.

The Xfer Entry cache depth will depend on the QA Chip Logical Interface implementation. For some implementations a single Xfer Entry value will be saved. If the Parameter Upgrader QA Device has no powersafe storage of Xfer Entry cache, a power down will cause the erasure of the Xfer Entry cache and the Parameter Upgrader QA Device will not be able to rollback to a pre-power-down value.

A dataset in the Xfer Entry cache will consist of the following:

-   -   Information about the Printer QA Device:

-   a. ChipId of the device.

-   b. FieldNum of the M0 field (i.e what was being upgraded).     -   Information about the Parameter Upgrader QA Device:

-   a. FieldNum of the M0 field used to transfer the count-remaining     from.     -   Xfer State—indicating at which state the transfer sequence is.         This will consist of:

-   a. State definition which could be one of the following:—Xfer,     StartRollBack and deleted (completed).

-   b. The value of sequence data fields SEQ_(—)1 and SEQ_(—)2.

The Xfer Entry cache stores the FieldNum of the count-remaining field of the Parameter Upgrader QA Device.

28.3.4.1 Adding New Dataset

A new dataset is added to Xfer Entry cache by the Xfer function.

There are three methods which can be used to add new dataset to the Xfer Entry cache. The methods have been listed below in the order of their priority:

-   1. Replacing existing dataset in Xfer Entry cache with new dataset     based on ChipId and FieldNum of the Ink QA Device in the new     dataset. A matching ChipId and FieldNum could be found because a     previous transfer output corresponding to the dataset stored in the     Xfer Entry cache has been correctly received and processed by the     Parameter Upgrader QA Device, and a new transfer request for the     same Printer QA Device, same field, has come through to the     Parameter Upgrader QA Device. -   2. Replace existing dataset cache with new dataset based on the Xfer     State. If the Xfer State for a dataset indicates deleted (complete),     then such a dataset will not be used for any further functions, and     can be overwritten by a new dataset. -   3. Add new dataset to the end of the cache. This will automatically     delete the oldest dataset from the cache regardless of the Xfer     State.     28.4 Upgrading the Count-Remaining Field

This section is only applicable to the Parameter Upgrader QA Device.

The transfer of count-remaining is similar to transfer ink-remaining because both involve transferring of amounts. Therefore, this transfer uses the XferAmount function.

The XferAmount function performs additional checks when transferring count-remaining. This includes checking of the operating parameter field, associated with the count-remaining. They are as follows:

-   -   The operating parameter value of the upgrading QA Device and the         QA Device being upgraded must match.     -   The operating parameter field (in both devices) must be         upgradeable by one key only, and all other keys must have         ReadOnly access. This key which has authenticated ReadWrite         permission to the operating parameter field, must be different         to the key that has authenticated Read Write permission to the         count-remaining field.     -   The data Type for the operating parameter field in the upgrading         QA Device must match the data Type for the operating parameter         field in the QA Device being upgraded.         28.5 New Operating Parameter Field Information         This section is only applicable to the Parameter Upgrader QA         Device.

This field stores the operating parameter value that is copied from the Parameter Upgrader QA Device to the operating parameter field being updated in the Printer QA Device.

This field has a single key associated with it. This key has authenticated ReadWrite permission to this field and will be referred to as write-parameter key.

Table 299 shows the field information for the new operating parameter field in the Parameter Upgrade QA Device

Attribute Name Value Explanation Type For e.g - Type describing the upgrade. TYPE_UPGRADE_PRINTSPEED_15^(a) KeyNum Slot number of the write-parameter Only the write-parameter key has authenticated key. ReadWrite access to this field. Non Auth 0 Non authenticated ReadWrite is not allowed RW Perm^(b) to the field. Auth RW 1 Authenticated (key based) ReadWrite access Perm^(c) is allowed to the field. KeyPerm KeyPerms[KeyNum] = 0 KeyNum is the slot number of the write-parameter key which has ReadWrite permission to the field. KeyPerms[others = 0 . . . 7] = 0 All other keys have ReadOnly access. End Pos Set as required. ^(a)This is a sample type only and is not included in the Type Map in Appendix A. ^(b)Non authenticated Read Write permission. ^(c)Authenticated Read Write permission. 28.6 Different Types of Transfer

There can be three types of transfer:

-   -   Parameter Transfer—This is transfer of an operating parameter         value from a Parameter Upgrader QA Device to a Printer QA         Device. This is performed when an upgradeable operating         parameter is written (for the first time) or changed.     -   Hierarchical refill—This is a transfer of count-remaining value         from one Parameter Upgrader Refill QA Device to a Parameter         Upgrader QA Device, where both QA Devices belong to the same         OEM. This is typically performed when OEM divides the number of         upgrades from one of its Parameter Upgrader QA Device to many of         its Parameter Upgrader QA Devices.     -   Peer to Peer refill—This is a transfer of count-remaining value         from one Parameter Upgrader Refill QA Device to Parameter         Upgrader Refill QA Device, where the QA Devices belong to         different organisations, say ComCo and OEM. This is typically         performed when ComCo divides number of upgrades from its         Parameter Upgrader QA Device to several Parameter Upgrader QA         Device belonging to several OEMs.

Transfer of count-remaining between peers, and hierarchical transfer of count-remaining, is similar to an ink transfer, but additional checks on the transfer request is performed when transferring count-remaining amounts. This is described in Section 28.4.1.

Transfer of an operating parameter value decrements the count-remaining by 1, hence is different to a ink-transfer.

FIG. 385 is a representation of various authorised upgrade paths in the printing system.

28.6.1 Hierarchical Transfers

Referring to FIG. 385, this transfer is typically performed when count-remaining amount is transferred from ComCo's Parameter Upgrader Refill QA Device to OEM's Parameter Upgrader Refill QA Device, or from QACo's Parameter Upgrader Refill QA Device to ComCo's Parameter Upgrader Refill QA Device.

This transfers are made using the XferAmount function (and not with the XferField described in Section 29.1). because count-remaining transfer is similar to fill/refilling of ink amounts, where ink amount is replaced by count-remaining amount.

28.6.1.1 Keys and Access Permission

We will explain this using a transfer from ComCo to OEM.

There is a count-remaining field associated with the ComCo's Parameter Upgrader Refill QA Device. This count-remaining field has two keys associated with:

-   -   The first key transfers count-remaining to the device from         another Parameter Upgrader Refill QA device(device is higher in         the heirachy), fills/refills the device itself.     -   The second key transfers count-remaining from it to other         devices (which are lower in the heirachy), fills/refills other         devices from it.

There is a count-remaining field associated with the OEM's Parameter Upgrader Refill QA Device.

-   -   This count-remaining field has a single key associated with:     -   This key transfers count-remaining to the device from another         Parameter Upgrader Refill QA device (which is higher or at the         same level in the heirachy), fills/refills (upgrades) the device         itself, and additionally transfers count-remaining from it to         other devices (which are lower in the heirachy), fills/refills         (upgrades) other devices from it.

For a successful transfer of count-remaining from ComCo's refill device to an OEM's refill device, the ComCo's refill device and the OEM's refill device must share a common key or a variant key.

This key is fill/refill key with respect to the OEM's refill device and it is the transfer key with respect to the ComCo's refill device.

For a ComCo to successfully fill/refill its refill device from another refill device (which is higher in the heirachy possibly belonging to the QACo), the ComCo's refill device and the QACo's refill device must share a common key or a variant key. This key is fill/refill key with respect to the ComCo's refill device and it is the transfer key with respect to the QACo's refill device.

28.6.1.1.1 Count-Remaining Field Information

Table 300 shows the field information for an _(M0) field storing logical count-remaining amounts in the refill device, which has the ability to transfer down the heirachy.

Attribute Name Value Explanation Type TYPE_COUNT_REMAINING^(a) Type describes that the field is a count- remaining field. KeyNum Slot number of the Only the refill key refill key. has authenticated ReadWrite access to this field. Non Auth 0 Non authenticated RW Perm^(b) ReadWrite is not allowed to the field. Auth RW 1 Authenticated (key Perm^(c) based) ReadWrite access is allowed to the field. KeyPerm KeyPerms[KeyNum] = 0 KeyNum is the slot number of the refill key, which has ReadWrite permission to the field. KeyPerms[Slot Num of Transfer key can transfer key ] = 1 decrement the field. KeyPerms[others = All other keys have 0 . . . 7(except ReadOnly access. transfer key)] = 0 End Pos Set as required. Depends on the amount of logical ink the device can store and storage resolution - i.e in picolitres or in microlitres. ^(a)Refer to Type Map in Appendix A for exact value. ^(b)Non authenticated Read Write permission. ^(c)Authenticated Read Write permission. 28.6.2 Peer to Peer Transfer

Referring to FIG. 385, this transfer is typically performed when count-remaining amount is transferred from OEM's Parameter Upgrader Refill QA Device to another Parameter Device Refill QA Device belonging to the same OEM.

28.6.2.1 Keys and Access Permission

There is an count-remaining field associated with the refill device. This count-remaining field has a single key associated with:

-   -   This key transfers count-remaining amount to the device from         another refill device (which is higher or at the same level in         the heirachy), fills/refills (upgrades) the device itself, and         additionally transfers ink from it to other devices (which are         lower in the heirachy), fills/refills (upgrades) other devices         from it.

This key is referred to as the fill/refill key and is used for both fill/refill and transfer. Hence, this key has both ReadWrite and Decrement-Only permission to the count-remaining field in the refill device.

28.6.2.1.1 Count-Remaining Field Information

Table 301 shows the field information for an _(M0) field storing logical count-remaining amounts in the refill device with the ability to transfer between peers.

TABLE 301 Field information for ink-remaining field for refill devices transferring between peers Attribute Name Value Explanation Type TYPE_COUNT_REMAINING^(a) Type describes that the field is a count- remaining field. KeyNum Slot number of the Only the refill key refill key. has authenticated ReadWrite access to this field. Non Auth 0 Non authenticated RW Perm^(b) ReadWrite is not allowed to the field. Auth RW 1 Authenticated (key Perm^(c) based) ReadWrite access is allowed to the field. KeyPerm KeyPerms[KeyNum] = 1 KeyNum is the slot number of the refill key, which has ReadWrite and Decre- ment permission to the field. KeyPerms[others = All other keys have 0 . . . 7(except ReadOnly access. KeyNum)] = 0 End Pos Set as required. Depends on the amount of logical ink the device can store and storage resolution - i.e in picolitres or in microlitres. ^(a)Refer to Type Map in Appendix A for exact value. ^(b)Non authenticated Read Write permission. ^(c)Authenticated Read Write permission. 29 Functions 29.1 XferField

-   -   Input KeyRef, _(M0)OfExternal, _(M1)OfExternal, ChipID,         FieldNumL, FieldNumE, InputParameterCheck (Optional), R_(E),         SIG_(E), R_(E2)     -   Output: ResultFlag, Field data, R_(L2), SIG_(out)     -   Changes: _(M0) and R_(L)     -   Availablity: Parameter Upgrader QA Device         29.1.1 Function Description

The XferField is similar to the XferAmount function in that it produces data and signature for updating a given _(M0) field. This data and signature when applied to the appropriate device through the WriteFieldsAuth function, will upgrade the FieldNumE (_(M0) field) of a device to the same value as FieldNumL of the upgrading device.

The system calls the XferField function on the upgrade device with a certain FieldNumL to be transferred to the device being upgraded The FieldNumE is validated by the XferField function according to various rules as described in Section 29.1.4. If validation succeeds the XferField function produces the data and signature for subsequent passing into the WriteFieldsAuth function for the device being upgraded.

The transfer field output consists of the new data for the field being upgraded, field data of the two sequence fields, and a signature. When a transfer output is produced, the sequence field data in SEQ_(—)1 is decremented by 2 from the previous value (as passed in with the input), and the sequence field data in SEQ_(—)2 is decremented by 1 from the previous value (as passed in with the input).

Additional InputParameterCheck value must be provided for the parameters not included in the SIG_(E), if the transmission between the System and Parameter Upgrader QA Device is error prone, and these errors are not corrected by the transimission protocol itself. InputParameterCheck is SHA-1[FieldNumL|FieldNumE|XferValLength|XferVal], and is required to ensure the integrity of these parameters, when these inputs are received by the Parameter Upgrader QA Device.

The XferField function must first calculate the SHA-1[FieldNumL|FieldNumE], compare the calculated value to the value received (InputParameterCheck) and only if the values match act upon the inputs.

29.1.2 Input Parameters

Table 302 describes each of the input parameters for XferField function.

Parameter Description KeyRef For common key input and output signature: KeyRef.keyNum = Slot number of the key to be used for testing input signature and producing the output signature. SIG_(E) produced using K_(KeyRef.keyNum) by the QA Device being upgraded. SIGout produced using K_(KeyRef.keyNum) for delivery to the QA Device being upgraded. KeyRef.useChipId = 0 For variant key input and output signatures: KeyRef.keyNum = Slot number of the key to be used for generating the variant key. SIG_(E) produced using a variant of K_(KeyRef.keyNum) by the QA Device being upgraded. SIGout produced using a variant of K_(KeyRef.keyNum) for delivery to the QA Device being upgraded. KeyRef.useChipId = 1 KeyRef.chipId = ChipId of the device which generated SIG_(E) and will receive SIGout. _(M0)OfExternal All 16 words of _(M0) of the QA Device being upgraded _(M1)OfExternal All 16 words of _(M1) of the QA Device being upgraded. ChipId ChipId of the QA Device being upgraded. FieldNumL _(M0) field number of the local (updating) device. The data stored in this field will be copied from the upgrading device. FieldNumE _(M0) field number of the QA Device being upgraded. This field will be updated to the value stored in FieldNumL within the upgrading device. R_(E) External random value used to verify input signature. This will be the R from the input signature generator (i.e device generating SIG_(E)). The input signal generator in this case, is the device being upgraded or a translation device. R_(E2) External random value used to produce output signature. This will be the R obtained by calling the Random function on the device which will receive the SIG_(out) from the XferField function. The device receiving the SIG_(out) in this case, is the device being upgraded or a translation device. SIG_(E) External signature required for authenticating input data. The input data in this case, is the output from the Read function performed on the device being upgraded. A correct SIG_(E) = SIG_(KeyRef)(Data | R_(E) | R_(L)). 29.1.2.1 Input Signature Verification Data Format

Refer to Section 27.1.2.1.

29.1.3 Output Parameters

Table 303 describes each of the output parameters for XferField function.

Parameter Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1, Table 292 and Table 303. FieldSelect Selection of fields to be written In this case the bit corresponding to SEQ_1, SEQ_2 and to FieldNumE are set to 1. All other bits are set to 0. FieldVal Updated data words for sequence data field and FieldNumE for QA Device being upgraded. Starts with LSW of lower field. This must be passed as input to the WriteFieldsAuth function of the QA Device being upgraded. R_(L2) Internal random value required to generate output signature This must be passed as input to the WriteFieldsAuth function or Translate function of the QA Device being upgraded. SIG_(out) Output signature which must be passed as an input to the WriteFieldsAuth function or Translate function of the QA Device being upgraded. SIG_(out) = SIG_(KeyRef)(data | R_(L2) | R_(E2)) as per FIG. 373

TABLE 303 Result Flag definitions for XferField ReultFlag Definition Description CountRemainingFieldInvalid The count- remaining field in Upgrading QA Device is invalid. FieldNumEKeyPermInvalid The upgrade field in the QA Device being upgraded doesn't have the correct authenticated permission. NoUpgradesRemaining The count-remaining field assocaited with the upgrade field in the Upgrading QA Device doesn't have any more upgrades left. 29.1.3.1 Output Signature Generation Data Format

Refer to Section 27.1.3.1.

29.1.4 Function Sequence

The XferField command is illustrated by the following pseudocode:

-   -   Accept input parameters-KeyRef, M0OfExternal, M1 OfExternal,         ChipID, FieldNumL, FieldNumE, R_(E), SIG_(E), R_(E2)

#Generate message for passing into ValidateKeyRefAndSignature function data

(RWSense|MSelect|KeyIdSelect|ChipId|WordSelect|M0|M1) # Refer to Figure 382. ---------------------------------------------------------------- # Validate KeyRef, and then verify signature ResultFlag = ValidateKeyRefAndSignature(KeyRef,data,R_(E),R_(L)) If (ResultFlag ≠ Pass) Output ResultFlag Return EndIf ---------------------------------------------------------------- # Validatate FieldNumE # FieldNumE is present in the device being upgraded PresentFlagFieldNumE

GetFieldPresent(M1OfExternal,FieldNumE) # Check FieldNumE present flag If(PresentFlagFieldNumE ≠ 1) ResultFlag

FieldNumEInvalid Output ResultFlag Return EndIf ---------------------------------------------------------------- # Check Seq fields exist and get their Field Number # Get Seqdata field SEQ_1 for the device being upgraded XferSEQ_1FieldNum

GetFieldNum(M1OfExternal, SEQ_1) # Check if the Seqdata field SEQ_1 is valid If(XferSEQ_1FieldNum invalid) ResultFlag

SeqFieldInvalid Output ResultFlag Return EndIf # Get Seqdata field SEQ_2 for the device being upgraded XferSEQ_2FieldNum

GetFieldNum(M1OfExternal, SEQ_2) # Check if the Seqdata field SEQ_2 is valid If(XferSEQ_2FieldNum invalid) ResultFlag

SeqFieldInvalid Output ResultFlag Return EndIf ------------------------------------------------------------------- --------------- #Check write permission for FieldNumE PermOKFieldNumE

CheckFieldNumEPerm(M1OfExternal, FieldNumE) If(PermOKFieldNumE ≠ 1) ResultFlag

FieldNumEWritePermInvalid Output ResultFlag Return EndIf ------------------------------------------------------------------- --------------- #Check that both SeqData fields have Decrement-Only permission with the same key #that has write permission on FieldNumE PermOKXferSeqData

CheckSeqDataFieldPerms(M1OfExternal, XferSEQ_1FieldNum, XferSEQ_2FieldNum,FieldNumE) If(PermOKXferSeqData ≠ 1) ResultFlag

SeqWritePermInvalid Output ResultFlag Return EndIf ------------------------------------------------------------------- ---------- # Get SeqData SEQ_1 data from device being upgraded GetFieldDataWords(XferSEQ_1FieldNum, XferSEQ_1DataFromDevice,M0OfExternal,M1OfExternal) # Get SeqData SEQ_2 data from device being upgraded GetFieldDataWords(XferSEQ_2FieldNum, XferSEQ_2DataFromDevice, M0OfExternal,M1OfExternal) ---------------------------------------------------------------- # FieldNumL(upgrade value)is a valid field in the upgrading device PresentFlagFieldNumL

GetFieldPresent(M1,FieldNumL) If(PresentFlagFieldNumL ≠ 1) ResultFlag

FieldNumLInvalid Output ResultFlag Return EndIf ---------------------------------------------------------------- #Get the CountRemaining field associated with the upgrade value field # The CountRemaining field is the next higher field from the upgrade value field FieldNumCountRemaining

FieldNumL + 1 # FieldNumCountRemaining is a valid field in the upgrading device PresentFlagFieldNumCountRemaining

GetFieldPresent(M1,FieldNumCountRemaining) If(PresentFlagFieldNumCountRemaining ≠ 1) ResultFlag

CountRemainingFieldInvalid Output ResultFlag Return EndIf ---------------------------------------------------------------- #Check permission for upgrade value field. Only one key (different # from KeRef.keyNum) has write permissions to the field and no key has decrement permissions. CheckOK

CheckUpgradeKeyForField(FieldNumL,M1,KeyRef) If(CheckOK ≠ 1) ResultFlag

FieldNumEKeyPermInvalid Output ResultFlag Return EndIf ---------------------------------------------------------------- #Find the type attribute for FieldNumE TypeFieldNumE

FindFieldNumType(M1OfExternal,FieldNumE) #Find the type attribute for FieldNumL (upgrade value) TypeFieldNumL

FindFieldNumType(M1,FieldNumL) If(TypeFieldNumE ≠ TypeFieldNumL) ResultFlag

TypeMismatch Output ResultFlag Return EndIf ---------------------------------------------------------------- # Check permissions for CountRemaining field # Check upgrades are available in the CountRemaining field of the # upgrading device i.e value of CountRemaining is non-zero positive number CountRemainingOK

CheckCountRemaining(FieldNumCountRemaining, M0, M1) If(CountRemainingOK ≠ 1) ResultFlag

NoUpgradesRemaining Output ResultFlag Return EndIf ---------------------------------------------------------------- #Get the size of the FieldNumL (upgrade value) If(FieldNumL = 0) FieldSizeOfFieldNumL

MaxWordInM− M1[FieldNumL].EndPos Else FieldSizeOfFieldNumL

M1[FieldNumL−1].EndPos− M1[FieldNumL].EndPos EndIf #Get the size of the FieldNumE (field being updated) If(FieldNumL = 0) FieldSizeOfFieldNumE

MaxWordInM− M1OfExternal[FieldNumE − 1].EndPos Else FieldSizeOfFieldNumE

M1OfExternal[FieldNumE−1].EndPos − M1OfExternal[FieldNumL].EndPos EndIf # Check whether the device being upgraded can hold the upgrade value from # FieldNumL If(FieldSizeOfFieldNumE < FieldSizeOfFieldNumL) ResultFlag

FieldNumESizeInsufficient Output ResultFlag Return EndIf ---------------------------------------------------------------- # All checks complete ..... # Generate Seqdata for SEQ_1 and SEQ_2 fields XferSEQ_1DataToDevice = XferSEQ_1DataFromDevice − 2 XferSEQ_2DataToDevice = XferSEQ_2DataFromDevice − 1 # Add DataSet to Xfer Entry Cache AddDataSetToXferEntryCache(ChipId,FieldNumE, FieldNumL, XferSEQ_1DataFromDevice, XferSEQ_2DataFromDevice) #Decrement CountRemaining field by one DecrementField(FieldNumCountRemaining,M0) #Get the upgrade value words from FieldNumE of the upgrading device GetFieldDataWords(FieldNumL,UpgradeValue,M0,M1) #Generate new field data words for FieldNumE. The upgrade value is copied to FieldDataE FieldDataE

UpgradeValue # Generate FieldSelect and FieldVal for SeqData field SEQ_1, SEQ_2 and # FieldDataE... CurrentFieldSelect

0 FieldVal

0 GenerateFieldSelectAndFieldVal(FieldNumE, FieldDataE, XferSEQ_1FieldNum, XferSEQ_1DataToDevice, XferSEQ_2FieldNum, XferSEQ_2DataToDevice, FieldSelect,FieldVal) #Generate message for passing into GenerateSignature function data

(RWSense|FieldSelect|ChipId|FieldVal)# Refer to Figure 373. #Create output signature for FieldNumE SIG_(out)

GenerateSignature(KeyRef,data,R_(L2),R_(E2)) Update R_(L2) to R_(L3) ResultFlag

Pass Output ResultFlag, FieldSelect,FieldVal, R_(L2) ,SIG_(out) Return EndIf 29.1.4.1 CountRemainingOK

CheckCountRemainingFieldNumL(FieldNumCountRemaining, M1, M0)

This functions checks permissions for CountRemaining field and also checks that upgrades are available in the CountRemaining field of the upgrading device.

AuthRW

M1[FieldNumCountRemaining].AuthRW NonAuthRW

M1[FieldNumCountRemaining].NonAuthRW DOForKeys

_(M1)[FieldNumCountRemaining].DOForKeys[KeyNum] Type

_(M1)[FieldNumCountRemaining].Type

TYPE_COUNT_REMAINING) PermOK

1 Else PermOK

0 Return PermOK EndIf #Get the count-remaining value from the upgrading device GetFieldDataWords(FieldNumCountRemaining,CountRemainingValue,M0,M1 ) If(CountRemainingValue <= 0) PermOK

0 Return PermOK EndIf PermOK

1 Return PermOK 29.2 RollBackField

-   -   Input: KeyRef, _(M0)OfExternal, _(M1)OfExternal, ChipID,         FieldNumL, FieldNumE, InputParameterCheck (optional), R_(E),         SIG_(E)     -   Output: ResultFlag     -   Changes: _(M0) and R_(L)     -   Availablity: Parameter Upgrader QA Device         29.2.1 Function Description

The RollBackField function is very similar to the RollBackAmount function, the only difference being that the RollBackField function adjusts the value of the count-remaining field associated with the upgrade value field of the upgrading device, instead of the upgrade value field itself. A successful rollback, increments the count-remaining by 1.

The Parameter Upgrader QA Device checks that the Printer QA Device didn't actually receive the transfer message correctly, by comparing the sequence data field values read from the device with the values stored in the Xfer Entry cache. The sequence data field values read must match what was previously written using the StartRollBack function. After all checks are fulfilled, the Parameter Upgrader QA Device adjusts its FieldNumL.

Additional InputParameterCheck value must be provided for the parameters not included in the SIG_(E), if the transmission between the System and Parameter Upgrader QA Device is error prone, and these errors are not corrected by the transimission protocol itself. InputParameterCheck is SHA-1[FieldNumL|FieldNumE], and is required to ensure the integrity of these parameters, when these inputs are received by the Parameter Upgrader QA Device.

The RollBackField function must first calculate the SHA-1[FieldNumL|FieldNumE], compare the calculated value to the value received (InputParameterCheck) and only if the values match act upon the inputs.

29.2.2 Input Parameters

Table 305 describes each of the input parameters for RollBackField function.

Parameter Description KeyRef For common key input signature: KeyRef.keyNum = Slot number of the key to be used for testing input signature. SIG_(E) produced using K_(KeyRef.keyNum) by the QA Device being upgraded. KeyRef.useChipId = 0 For variant key input signature: KeyRef.keyNum = Slot number of the key to be used for generating the variant key. SIG_(E) produced using a variant of K_(KeyRef.keyNum) by the QA Device being upgraded. KeyRef.useChipId = 1 KeyRef.chipId = ChipId of the device which generated SIG_(E.) _(M0)OfExternal 16 words of _(M0) of the QA Device being upgraded which failed to upgrade. _(M1)OfExternal 16 words of _(M1) of the QA Device being upgraded which failed to upgrade. ChipId ChipId of the QA Device being upgraded which failed to upgrade. FieldNumL _(M0) field number of the local (upgrading) device whose value could not be copied to the device being upgraded. FieldNumE _(M0) field number of the QA Device being upgraded to which the upgrade value in FieldNumL couldn't be copied. R_(E) External random value used to verify input signature. This will be the R from the input signature generator (i.e device generating SIG_(E)). The input signal generator in this case, is the device which failed to upgrade or a translation device. SIG_(E) External signature required for authentic- ating input data. The input data in this case, is the output from the Read function performed on the device which failed to upgrade. A correct SIG_(E) = SIG_(KeyRef)(Data | R_(E) | R_(L)). 29.2.2.1 Input Signature Generation Data Format

Refer to Section 27.1.2.1 for details.

29.2.3 Output Parameters

Table 306 describes each of the output parameters for RollBackField.

Parameter Description ResultFlag Indicates whether the function completed successfully or not. If it did not complete successfully, the reason for the failure is returned here. See Section 12.1, Table 292, Table 304 and Table 295.

29.2.4 Function Sequence

The RollBackField command is illustrated by the following pseudocode:

Accept input parameters-KeyRef, M0OfExternal, M1OfExternal, ChipId, FieldNumL, FieldNumE, R_(E),SIG_(E) #Generate message for passing into GenerateSignature function data

(RWSense|MSelect|KeyIdSelect|ChipId|WordSelect|M0|M1) # Refer to Figure 382. ---------------------------------------------------------------- # Validate KeyRef, and then verify signature ResultFlag = ValidateKeyRefAndSignature(KeyRef,data,R_(E),R_(L)) If (ResultFlag ≠ Pass) Output ResultFlag Return EndIf ---------------------------------------------------------------- # Check Seq fields exist and get their Field Number # Get Seqdata field SEQ_1 num for the device being upgraded XferSEQ_1FieldNum

GetFieldNum(M1OfExternal, SEQ_1) # Check if the Seqdata field SEQ_1 is valid If(XferSEQ_1FieldNum invalid) ResultFlag

SeqFieldInvalid Output ResultFlag Return EndIf # Get Seqdata field SEQ_2 num for the device being upgraded XferSEQ_2FieldNum

GetFieldNum(M1OfExternal, SEQ_2) # Check if the Seqdata field SEQ_2 is valid If(XferSEQ_2FieldNum invalid) ResultFlag

SeqFieldInvalid Output ResultFlag Return EndIf ---------------------------------------------------------------- # Get SeqData SEQ_1 data from device being upgraded GetFieldDataWords(XferSEQ_1FieldNum, XferSEQ_1DataFromDevice,M0OfExternal,M1OfExternal) # Get SeqData SEQ_2 data from device being upgraded GetFieldDataWords(XferSEQ_2FieldNum, XferSEQ_2DataFromDevice, M0OfExternal,M1OfExternal) # Generate Seqdata for SEQ_1 and SEQ_2 fields with the data that is read XferSEQ_1Data = XferSEQ_1DataFromDevice + 1 XferSEQ_2Data = XferSEQ_2DataFromDevice + 2 # Check Xfer Entry in cache is correct - dataset exists, Field data # and sequence field data matches and Xfer State is correct XferEntryOK

CheckEntry(ChipId, FieldNumE, FieldNumL, XferSEQ_1Data, XferSEQ_2Data) If( XferEntryOK= 0) ResultFlag

RollBackInvalid Output ResultFlag Return EndIf # Increment associated CountRemaining by 1 IncrementCountRemaining(FieldNumCountRemaining) # Update XferState in DataSet to complete/deleted UpdateXferStateToComplete(ChipId,FieldNumE) ResultFlag

Pass Output ResultFlag Return Example Sequence of Operations 30 Concepts

The QA Chip Logical Interface interface devices do not initiate any activities themselves. Instead the System reads data and signature from various untrusted devices, and sends the data and signature to a trusted device for validation of signature, and then uses the data to perform operations required for printing, refilling, upgrading and key replacement. The system will therefore be responsible for performing the functional sequences required for printing, refilling, upgrading and key replacement. It formats all input parameters required for a particular function, then calls the function with the input parameters on the appropriate QA Chip Logical Interface instance, and then processes/stores the output parameters from the function appropriately.

Validation of signatures is achieved by either of the following schemes:

-   -   Direct—the signature produced by an untrusted device is directly         passed in for validation to the trusted device. The direct         validation requires the untrusted device to share a common key         or a variant key with the trusted device. Refer to Section 7 for         further details on common and variant keys.     -   Translation—the signature produced by an untrusted is first         validated by the translating device, and a new signature of the         read data is produced by the translation device for validation         by the trusted device. Several translation device may be chained         together—the first translation device validates the signature         from the untrusted device, and the last translation device         produces the final signature for validation by the trusted         device. The translation device must share a common key or a         variant key with the trusted/untrusted device and among         themselves, if several translation devices are chained together         for signature validation.         30.1 Representation

Each functional sequence consists of the following devices (refer to Section 4.3):

-   -   System.     -   A trusted QA Device—which may be a system trusted QA Device, or         an Parameter Upgrader QA Device, or a Ink Refill QA Device, or a         Key Programmer QA Device depending on the function performed.         This device is referred to as device A.     -   An untrusted QA Device—which may be a Printer QA Device, or an         Ink QA Device. This device is referred to as device B.     -   A translation QA Device will be used if a translation scheme is         used to validate signatures. This device is referred to as         device C.

The command sequence produced by the system for further sequences will be documented as shown in Table 307.

TABLE 307 Command sequence representation Sequence No Function Parameters Sequence Device.FunctionName Input Parameters and their values. order Output parameters and their description.

Therefore, a typical direct signature validation sequence can be represented by FIG. 386 and Table 308.

For a direct signature to be used, A and B must share a common or a variant key i.e B.K_(n1)=A.K_(n2) or B.K_(n1)=FormKeyVariant(A.K_(n2), B.ChipID).

TABLE 308 Command sequence for direct signature validation Sequence No Function Parameters 1 A.Random None R_(A) = RL 2 B.Read KeyRef = n1, SigOnly = 0, MSelect = Any one M, KeyIdSelect = 0, WordSelectForDesiredM = Any one word in the selected M, RE = R_(A) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per input [MSelect] and [WordSelectForDesiredM], R_(B) = R_(L), SIG_(B) = SIGout Refer to Section 15.3.1. 3 A.Test KeyRef = n2, DataLength = Length of MWords in words preformatted as per Section 16.1, Data = MWords preformatted as per Section 16.1, RE = R_(B), SIGE = SIG_(B) ResultFlag = Pass/Fail

A typical signature validation using translation can be represented by FIG. 387 and Table 309.

For validating signatures using translation:

-   -   A and C must share a common or a variant key         -   i.e C.K_(n3)=A.K_(n2) or C.K_(n3)=FormKeyVariant(A.K₂,             C.ChipID).     -   B and C must share a common or a variant key         -   i.e C.K_(n2)=B.K_(n1) or B.K_(n1)=Form KeyVariant(C.K_(n2),             B.ChipID).

TABLE 309 Command sequence for signature validation using translation Sequence No Function Parameters 1 C.Random None R_(C) = RL 2 B.Read KeyRef = n1, SigOnly = 1 or 0, MSelect = any, KeyIdSelect = any, WordSelectForDesiredM = any, RE = R_(C) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per input [MSelect] and [WordSelectForDesiredM], R_(B) = R_(L), SIG_(B) = SIGout Refer to Section 15.3.1. 3 A.Random None R_(A) = RL 4 C.Translate InputKeyRef = n2, DataLength = Length of MWords in words preformatted as per Section 17.1, Data = MWords preformatted as per Section 17.1, RE = R_(B), SIGE = SIG_(B), OutputKeyRef = n3, RE2 = R_(A) If ResultFlag = Pass then R_(C1) = R_(L2), SIG_(C) = SIGOut Refer to Section 15.3.1 5 A.Test KeyRef = n2, DataLength = Length of MWords in words preformatted as per Section 16.1, Data = MWords preformatted as per Section 16.1, RE = R_(C1), SIGE = SIG_(C) ResultFlag = Pass/Fail 31 In Field Use

This section covers functional sequences for printer and ink QA Devices, as they perform their usual function of printing.

31.1 Startup Sequence

At startup of any operation (a printer startup or an upgrade startup), the system determines the properties of each QA Device it is going to communicate with. These properties are:

-   -   Software version of the QA Device. This includes         SoftwareReleaseIdMajor and SoftwareReleaseIdMinor. The         SoftwareReleaseIdMajor identifies the functions available in the         QA Device. Refer to Section 13.2 for details.     -   The number of memory vectors in the QA Device.     -   The number of keys in the QA Device.     -   The ChipId of the QA Device.

The properties allow the system to determine which functions are available in a given QA Device, as well as the value of input parameters required to communicate with the QA Device.

TABLE 310 Startup command sequence Sequence No Function Command 1 B.GetInfo None Major release identifier of the QA Device = SoftwareReleaseIdMajor, Minor release identifier of the QA Device = SoftwareReleaseIdMinor, Number of memory vectors in the QA Device = NumVectors, Number of keys in the QA Device = NumKeys, Id of the QA Device = ChipId 0 = VarDataLen No VarData in case of an ink or printer QA Device 31.1.1 Clearing the Preauthorisation Field

Preauthorisation of ink is one of the schemes that a printer may use to decrement logical ink as physical ink is used. This is discussed in details in Section 31.4.3.

If the printer uses preauthorisation, the system must read the preauthorisation field at startup. If the preauthorisation field is not clear, then the system must apply (decrement) the preauth amount to the corresponding ink field, by performing a non-authenticated write of the decremented amount to the appropriate ink field, and then clear the preauthorisation field by performing an authenticated write to the preauthorisation field.

31.2 Presence Only Authentication

The purpose of presence only authentication is to determine whether the printer should or shouldn't work with the ink cartridge.

31.2.1 Without Data Interpretation

This sequence is performed when the printer authenticates the ink cartridge. The authentication consists of verifying a signature generated by the untrusted ink QA Device (in the ink cartridge) using the system's trusted QA Device.

For signature to be valID, the trusted QA Device (A) and the untrusted ink QA Device (B) must share a common or a variant key i.e B.K_(n1)=A.K_(n2) or B.K_(n1)=FormKeyVariant(A.K_(n2), B.ChipID).

A single word of a single M is read because the system is only interested in the validity of signature for a given data.

If the printer wants to verify the signature and doesn't require any data from the ink cartridge (because it is cached in the printer), then the printer calls the Read function with SigOnly set to 1. The Read returns only the signature of the data as requested by the input parameters. The printer then sends its cached data and signature (from the Read function) to its trusted QA Device for verification. The printer may use this signature verification scheme if it has read the data previously from the ink QA Device, and the printer knows that the data in the ink QA Device has not changed from value that was read earlier by the printer.

Table 311 shows the command sequence for performing presence only authentication requiring both data and signature.

Seq No Function Parameters 1 A.Random None R_(A) = RL 2 B.Read KeyRef = n1, SigOnly = 0, MSelect = Any one M, KeyIdSelect = 0, WordSelectForDesiredM = Any one word in the selected M, RE = R_(A) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per input [MSelect] and [WordSelectForDesiredM], R_(B) = R_(L), SIG_(B) = SIGout Refer to Section 15.3.1. 3 A.Test KeyRef = n2, DataLength = Length of MWords in words preformatted as per Section 16.1, Data = MWords preformatted as per Section 16.1, RE = R_(B), SIGE = SIG_(B) ResultFlag = Pass/Fail 31.2.2 With Data Interpretation

This sequence is performed when the printer reads the relevant data from the untrusted QA Device in the ink cartridge. The system validates the signature from the external ink QA Device, and then uses this data for further processing.

For signature to be valID, the trusted QA Device (A) and the untrusted QA Device (B) must share a common or a variant key i.e B.K_(n1)=A.K_(n2) or B.K_(n1)=FormKeyVariant(A.K_(n2), B.ChipID).

The data read assists the printer to determine the following before printing can commence:

-   -   Which fields in _(M0) store logical ink amounts in the ink QA         Device.     -   The size of the ink fields in the ink QA Device. Refer to         Section 8.1.1.1.     -   The type of ink.     -   The amount of ink in the field.

Table 312 shows the command sequence for performing presence only authentication (with data interpretation).

Seq No Function Parameters 1 A.Random None R_(A) = RL 2 B.Read KeyRef = n1, SigOnly = 0, MSelect = 0x03(indicates M0 and M1), KeyIdSelect = 0xFF (Read all KeyIds), WordSelectForDesiredM (for _(M0)) = 0xFFFF (Read all 16 _(M0)words), WordSelectForDesiredM (for_(M1)) = 0xFFFF(Read all 16_(M1)words), RE = R_(A) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per input [MSelect] and [WordSelectForDesiredM], All 16 words of _(M0) and _(M1). R_(B) = RL SIG_(B) = SIGout Refer to Section 15.3.1 3 A.Test Input Key = n2, DataLength = Length of MWords in words preformatted as per Section 16.1, Data = MWords preformatted as per Section 16.1, RE = R_(B), SIGE = SIG_(B) ResultFlag = Pass/Fail 31.2.2.1 Locating Ink Fields and Determining Ink Amounts Remaining

Before printing can commence, the printer must determine the ink fields in the ink cartridge so that it can decrement these fields with the physical use of ink. The printer must also verify that the ink in the ink cartridge is suitable for use by the printer.

This process requires reading data from the ink QA Device and then comparing the data to what is required. To perform the comparison the printer must store a list for each ink it uses.

The ink list must consist of the following:

-   -   Ink Id—A identifier for the ink     -   KeyId—The KeyId of the key used to fill/refill this ink.     -   Type—This is the type attribute of the ink.

The ink list stored in the printer is shown in Table 313.

Ink Id KeyId Type 1- represents 1- represents KeyId of 0x55 black ink Network_OEM_InkFill/Refill Key^(b) TYPE_REGULAR_BLACK_INK^(a) 2- represents 1- represents KeyId of 0x9F cyan ink Network_OEM_InkFill/Refill Key^(b) TYPE_HIGHQUALITY_CYAN_INK^(a) 3- represents 1- represents KeyId of 0x9A magenta ink Network_OEM_InkFill/Refill Key^(b) TYPE_HIGHQUALITY_MAGENTA_INK^(a) 4- represents 1- represents KeyId of 0x9C yellow ink Network_OEM_InkFill/Refill Key^(b) TYPE_HIGHQUALITY_YELLOW_INK^(a) ^(a)These Types are only used as an example. ^(b)These KeyIds are only used as an example.

-   -   a. These Types are only used as an example.     -   b. These Keyids are only used as an example.

The printer will perform a Read of the ink QA Device's M0, M1 and KeyIds to determine the following:

-   -   The correct ink field (_(M0) field) in the ink QA Device.     -   The amount of ink-remaining in the field.

The ink QA Device's M1 and KeyId helps the printer determine the location of the ink field and ink QA Device's M0 and M1 helps determine the amount of ink-remaining in the field.

31.2.2.2 FieldNum FindFieldNum(keyIdRequired, typeRequired)

This function returns a FieldNum of an M0 field, whose authenticated ReadWrite access key's KeyId is keyIdRequired, and whose Type attribute matches typeRequired. If no matching field is found it returns a FieldNum=255. This function must be available in the printer system so that it can determine the ink field required by it.

The function sequence is described below.

# Get total number of fields in the ink QA Device FieldSize[16]

0 # Array to hold FieldSize assuming there are 16 fields NumFields

FindNumberOfFieldsInM0(M1,FieldSize) # Refer to Section 19.4.1. # Loop through KeyIds read assuming all KeyIds have been read from ink QA Device For i

0 to 7 #Check if KeyId read matches If(KeyId_(i) = keyIdRequired # Matching KeyId found KeyNum

i # Get the KeyNum of the matching KeyId # Now look through the field to check which field has #write permissions with this KeyNum For j

0 to NumOfFields AuthRW

_(M1)[j].AuthRW # Isolate AuthRW for field # Check authenticated write is allowed to the field If(AuthRW = 1) KeyNum_(j)

_(M1)[j].KeyNum # Isolate KeyNum of the field Typej

_(M1)[j].Type #Islotate Type attribute of the field # Check if Key is write key for the field and type of Ink Id#2 If (KeyNum = KeyNum_(j))

(Type_(j) = typeRequired) FieldNum

j return FieldNum EndIf EndIf EndFor # Loop through to next field FieldNum

255 # Error - no field found return FieldNum EndIf EndFor # Loop through to next KeyId

-   -   For e.g if the printer wants to find an ink field that matches         Ink Id#2 (from Table 313) in the ink QA Device, it must call the         function FindFieldNum with keyIdRequired=Keyld of         Network_OEM_InkFill/Refill Key and         typeRequired=TYPE_HIGHQUALITY_CYAN_INK.         31.2.2.3 Ink-Remaining Amount

This can be determined by using the function GetFieldDataWords(FieldNum,FieldData

, M0,M1) described in Section 27.1.4.14. FieldNum must be set to the value returned from function in Section 31.2.2.2. FieldData returns the ink-remaining amount.

The function GetFieldDataWords(FieldNum,FieldData

, M0,M1) must be implemented in the printer system.

31.3 Presence Only Authentication Through the Translate Function

This sequence is performed when the printer reads the data from the untrusted ink QA Device in the ink cartridge but uses a translating QA Device to indirectly validate the read data. The translating QA Device validates the signature using the key it shares with the untrusted QA Device, and then signs the data using the key it shares with the trusted QA Device. The trusted QA Device then validates the signature produced by the translating QA Device.

For validating signatures using translation:

-   -   A and C must share a common or a variant key         -   i.e C.K_(n3)=A.K₂ or C.K_(n3)=FormKeyVariant(A.K_(n2),             C.ChipID).     -   B and C must share a common or a variant key         -   i.e C.K_(n2)=B.K_(n1) or B.K_(n1)=FormKeyVariant(C.K_(n2),             B.ChipID).

Table 314 shows a command sequence for presence only authentication using translation

Seq No Function Parameters 1 C.Random None R_(C) = RL 2 B.Read KeyRef = n1, SigOnly = 1 or 0, MSelect = any M, KeyIdSelect = 0, WordSelectForDesiredM = any, RE = R_(C) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per input [MSelect] and [WordSelectForDesiredM], R_(B) = R_(L), SIG_(B) = SIGout Refer to Section 15.3.1 3 A.Random None R_(A) = RL 4 C.Translate InputKeyRef = n2, DataLength = Length of MWords in words preformatted as per Section 17.1, Data = MWords preformatted as per Section 17.1, RE = R_(B), SIGE = SIG_(B), OutputKeyRef = n3, RE2 = R_(A) If ResultFlag = Pass then R_(C1) = RL1, SIG_(C) = SIGOut Refer to Section 15.3.1 5 A.Test KeyRef = n2, DataLength = Length of MWords in words preformatted as per Section 16.1, Data = MWords preformatted as per Section 16.1, RE = R_(C1), SIGE = SIG_(C) ResultFlag = Pass/Fail 31.4 Updating the Ink-Remaining

This sequence is performed when the printer is printing. The ink QA Device holds the logical amount of ink-remaining corresponding to the physical ink left in the cartridge. This logical ink amount must decrease, as physical ink from the ink cartridge is used for printing.

31.4.1 Sequence of Update

The primary question is when to deduct the logical ink amount—before or after the physical ink is used.

-   a. Print first (use physical ink) and then update the logical ink.     If the power is cut off after a physical print and before a logical     update, then the logical update is not performed. Therefore, the     logical ink-remaining is more than the physical ink-remaining.     Performing repeated power cuts will increase the differential     amount, and finally any physical ink could be used to refill the QA     Device. -   b. Update the logical ink and then print (use physical ink). This is     better than (a) because other physical inks cannot be used. However,     if a problem occurs during printing, after the logical amount has     already been deducted, there will be a disparity between logical and     physical amounts. This might result in the printer not printing even     if physical ink is present in the ink cartridge. The amount of     disparity can be reduced by increasing the frequency of updating     logical ink i.e update after each line instead of after each page. -   c. Preauthorise logical ink. Preauthorise certain amount of ink     (depends on the frequency of logical updates) before print and clear     it at the end of printing. If power is cut off after a page is     printed, then on start up, the printer reads the preauthorisation     field, if it has not been cleared, it applies the preauth amount to     the ink-remaining amount, and then clears the preauthorisation     field.     31.4.2 Basic Update

Some printers may use one of methods described in Section 31.4.1 (a) or (b) to update logical ink amounts in the ink QA Device. This method of updating the ink is termed as a basic update. The decremented amount is written to the appropriate ink field (which has been previously determined using Section 31.2.2) in _(M0) The printer verifies the write, by reading the signature of the written data, then passing it to the Test function of the trusted QA Device.

For signature to be valID, the trusted QA Device (A) and ink QA Device (B) must share a common or a variant key i.e B.K_(n1)=A.K_(n2) or B.K_(n1)=FormKeyVariant(A.K_(n2), B.ChipID).

TABLE 315 Command sequence for updating the ink-remaining (basic) Seq No Function Parameter 1 B.WriteFields VectNum = 0, FieldSelect = Select bits corresponding to the Ink fields, The ink field locations should have been determined before by using the method in Section ‘31.2.2.1 FieldVal = Decremented ink-remaining amount ResultFlag = Pass/Fail 2 A.Random None R_(A) = RL 3 B.Read KeyRef = n1, SigOnly = 1, (We only need the signature because we already know the data) MSelect = _(M0), KeyIdSelect = 0, WordSelectForDesiredM = corresponds to the ink fields written in Seq No 1, RE = R_(A) If ResultFlag = Pass then SelectedWordsOfSelectedMs not returned because [SigOnly] = 1 in Seq 3, R_(B) = R_(L), SIG_(B) = SIGout Refer to Section 15.3.1 4 A.Test KeyRef = n2, DataLength = length in words as per Seq No 1 [MVal] preformatted as per Section 16.1, Data = as per Seq No 1 [MVal] preformatted as per Section 16.1, RE = R_(B), SIGE = SIG_(B) ResultFlag = Pass/Fail 31.4.3 Preauthorisation

This section describes the update of logical ink amounts using preauthorisation.

The basic preauthorisation sequence is as follows:

-   a. Preauthorise before the first print. Preauthorisation amount     depends on the printer model. Example amounts could be the ink     required for an fully covered A4 page or an A3 page. Value     corresponding to the preauth amount is written to the preauth field     in the ink QA Device.

Note: The preauth value must be correctly interpreted on different printer models i.e if a preauthorisation amount of A4 page is set in the ink cartridge in printer1(model1), and later the ink cartridge is placed in printer2(model2) with its preauth still set, printer2 must deduct an A4 page worth of ink from ink-remaining amount.

-   b. Print the page. -   c. Write the deducted logical amount to the ink field of the ink QA     Device and validate the write by reading the signature of the ink     field. -   d. Repeat b to c till the last page has been printed. -   e. Clear the preauth amount. -   f. If the power is cut off before the preauth is applied, on startup     apply the preauth amount to the corresponding ink field, by     performing a non authenticated write of the decremented amount and     clear the preauth amount by performing an authenticated write of the     preauth field.     31.4.3.1 Set Up of the Preauth Field

Only a single preauth gield must exist in an Ink QA Device.

Preauth field will consist of a single _(M0) word but can be optionally extended to two _(M0) words by using a different value of type attribute. FIG. 388 shows the setup of preauth field's attributes in _(M1).

The preauth field has authenticated ReadWrite access using the INK_USAGE_KEY i.e INK_USAGE_KEY can perform authenticated writes to this field. This key or its variant is shared between the ink QA Device and the printer QA Device to validate any data read from the ink cartridge. For signature to be valID, B.K_(n1)=A.K_(n2) or B.K_(n1)=FormKeyVariant(A.K_(n2), B.ChipID), where K_(n1)=INK_USAGE_KEY. The system performs a WriteAuth to the preauth field using this key, to set up the preauth amount, and to clear the preauth amount.

The preauth field is identified by two attributes:

-   -   Type attribute—TYPE_PREAUTH. Refer to Appendix A.     -   KeyId of KeyNum attribute must be the same as the KeyId of the         INK_USAGE_KEY which the printer uses to validate the any data         read from the ink QA Device.

The Preauth field can be applied to a single ink field or multiple ink fields.

31.4.3.2 Preauth Applied to a Single Ink Field

In this case the entire preauth field is used to store the preauth amount and is only linked to one ink field.

31.4.3.3 Preauth Applied to Multiple Ink Fields

Multiple preauth fields can be accommodated in a single M₀ field by a scheme shown in FIG. 388A.

This scheme supports a maximum of 8 ink fields being present in the Ink QA Device.

The field in _(M0) is divided into two parts— preauth field select and preauth amount. Each bit in preauth field select corresponds to a single ink field, and the preauth amount for each ink field is the same. If an ink cartridge uses multiple inks which are preauthorised, then each of the inks will have a corresponding preauth field bit. Before a particular ink is used for printing the corresponding preauth field bit is set. The preauth amount field is also set if the previous amount is zero. At finish, the preauth field bit is cleared. If more than one ink is used, the preauth bit for each ink field is set, and at finish each bit is cleared with last bit clearing the preauth amount as well.

31.4.3.4 Locating Preauth Fields and Determining Preauth Field Value

The preauth field can be located in the same manner as the ink field. If the printer wants to find the preauth field in the ink QA Device, it must call the function FindFieldNum (see Section 31.2.2.2) with keyIdRequired=KeyId of Network_OEM_Ink_Usage_Key and typeRequired=TYPE_PREAUTH. The preauth field value can be read in the same manner as the ink-remaining amount. This requires using of the function GetFieldDataWords(FieldNum,FieldData

, M0,M1) described in Section 27.1.4.14. FieldNum must be set to the value returned from function FindFieldNum, which in this case is the field number of the preauth field. FieldData returns the value of the preauth field.

31.4.3.5 Command Sequence

The command sequence can be broken up into three parts:

-   -   Start of print sequence.     -   During print sequence.     -   End of print sequence.         31.4.3.5.1 Start of Print Sequence

This sets up the preauth amount before the start of printing.

Table 316 shows the command sequence for start of print sequence. The first Random-Read-Test sequence determines the preauth field in the ink QA Device and its value. The Random-SignM-WriteFieldsAuth sequence, then writes to the preauth field the new preauth value.

TABLE 316 Updating the consumable remaining (preauth) start of print sequence Seq No Function Parameters Random-Read -Test sequence to determine the location of the preauth field in the ink QA Device and its value 1 A.Random None R_(A) = RL 2 B.Read KeyRef = n1, SigOnly = 0, WordSelectForDesiredM (for _(M0)) = all 16 words of M0 and all 16 words of M1 MSelect = 0x03(indicates M0 and M1), KeyIdSelect = 0xFF (Read all KeyIds), WordSelectForDesiredM (for _(M0)) = 0xFFFF (Read all 16 _(M0)words), WordSelectForDesiredM (for _(M1)) = 0xFFFF(Read all 16 _(M1)words), RE = R_(A) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per input [MSelect] and [WordSelectForDesiredM], R_(B) = R_(L), SIG_(B) = SIGout Refer to Section 15.3.1 3 A.Test KeyRef = n2, DataLength = length of MWords in words preformatted as per Section 16.1, Data = MWords as per Seq No 2 preformatted as per Section 16.1, RE = R_(B), SIGE = SIG_(B) ResultFlag = Pass/Fail Random-SignM-WriteFieldsAuth sequence to write the new preauth value 4 B.Random None R_(B1) = RL 5 A.SignM KeyRef = n2, FieldSelect = Select bit corresponding to the Preauth field, FieldVal = new preauth value, ChipId = ChipId of B, R_(E) = R_(B1) If ResultFlag = Pass then R_(A1) = R_(L) SIG_(A) = SIGout Refer to Section 27.1.3.1 6 B.WriteFieldsAuth KeyRef = n1, FieldSelect = same as Seq 5 [FieldSelect], FieldVal = same as Seq 5 [FieldVal], RE = R_(A1), SIGE = SIG_(A) ResultFlag = Pass/Fail 31.4.3.5.2 During Print Sequence

This set of commands are repeated at equal intervals to update logical ink amounts to the ink QA Device during printing.

Table 317 shows the command sequence for the print sequence. The WriteFields writes the updated value to the ink field. Random-Read-Test reads back the value written and tests whether the value read matches the value written.

TABLE 317 Updating the consumable remaining (preauth) during print sequence Seq No Function Parameters Write the decremented ink-remaining account. 7 B.WriteFields FieldSelect = Select bits corresponding to the Ink fields, FieldVal = Decremented ink-remaining amount for a single ink or multiple ink fields as per FieldSelect. ResultFlag = Pass/Fail Random-Read-Test sequence to read and verify the ink-remaining amount written 8 A.Random None R_(A) = RL 9 B.Read KeyRef = n1, SigOnly = 1 -(We only need the signature because we already know the data), MSelect = 0x01 (only _(M0)), KeyIdSelect = 0, WordSelectForDesiredM = corresponds to the ink fields written in Seq No 7, RE = R_(A) If ResultFlag = Pass then SelectedWordsOfSelectedMs not returned because [SigOnly] = 1 in Seq 9 R_(B) = R_(L), SIG_(B) = SIGout Refer to Section 15.3.1. 10 A.Test KeyRef = n2, DataLength = length in words as per Seq No 7 [MVal] preformatted as per Section 16.1, Data = as per Seq No 7 [MVal] preformatted as per Section 16.1, RE = R_(B), SIGE = SIG_(B) ResultFlag = Pass/Fail 31.4.3.5.3 End of Print Sequence

This sequence clears preauth amount before the print sequence is completed.

Table 318 shows the command sequence for the end of print sequence.

The preauth field is read using the Random-Read-Test sequence. And the preauth field is cleared using the Random-SignM-WriteFieldsAuth sequence.

TABLE 318 Updating the consumable remaining (preauth) end of print sequence Seq No Function Parameters Random-Read-Test sequence to read the preauth field and verify the preauth data 11 A.Random None R_(A) = R_(L) 12 B.Read KeyRef = n1, SigOnly = 1, MSelect = 0x01(only M0), KeyIdSelect = 0, WordSelectForDesiredM (for _(M0)) = Words corresponding to the Pre- authfield that has been written to in Seq 5 [FieldSelect] in Table 317. RE = R_(A) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per Seq No 12 [MSelect] and [WordSelectForDesiredM], R_(B) = R_(L), SIG_(B) = SIGout Refer to Section 15.3.1 13 A.Test KeyRef = n2, DataLength = length of MWords in words as per Seq No 12 preformatted as per Section 16.1, Data = MWords as per Seq No 12 preformatted as per Section 16.1, RE = R_(B), SIGE = SIG_(B) ResultFlag = Pass/Fail Random-SignM-WriteFieldsAuth sequence clears the preauth field 14 B.Random None R_(B1) = R_(L) 15 A.SignM KeyRef = n2, FieldSelect = Select bit corresponding to Pre authfield, FieldVal = Clear the preauth field, ChipId = ChipId of B, R_(E) = R_(B1) If ResultFlag = Pass then R_(A1) = R_(L) SIG_(A) = SIGout Refer to Section 27.1.3.1 16 B.WriteFieldsAuth KeyRef = n1, FieldNum = same as Seq 5 [FieldSelect], FieldData = same as Seq 5 [FieldVal], RE = R_(B1), SIGE = SIG_(A) ResultFlag = Pass/Fail 31.4.4 Preauthorisation Through the Translate Function

This is performed when the system trusted QA Device doesn't share a key with the ink QA Device, and uses a translating QA Device to Translate a Read from the ink QA Device, and to Translate a SignM to the ink QA Device.

The basic translate principle involves translating the Read data from the untrusted QA Device, to the Test data of the trusted QA Device, and translating the SignM data from the trusted QA Device, to the WriteFieldsAuth data of the untrusted QA Device.

For validating signatures using translation:

-   -   The trusted QA Device (A) and the translating QA Device (C) must         share a common or a variant key i.e C.K_(n3)=A.K_(n2) or         C.K₃=FormKeyVariant(A.K_(n2), C.ChipID).     -   The ink QA Device (B) and the translating QA Device (C) must         share a common or a variant key i.e C.K_(m2)=B.K_(n1) or         B.K_(n1)=FormKeyVariant(C.K_(n2), B.ChipID).

Only the start of print sequence is described using Translate. The rest of the sequences in preauthorisation can be modified to apply translation using this example.

Table 319 shows the command sequence for preauth (start of print sequence) using translation.

TABLE 319 Preauth(start of print sequence) using translate command Seq No Function Parameter Random-Read-Random-Translate-Test sequence reads the location of the preauth field and its value using the translating QA Device C 1 C.Random None R_(C) = RL 2 B.Read KeyRef = n1, SigOnly = 0, MSelect = 0x03(indicates M0 and M1), KeyIdSelect = 0xFF (Read all KeyIds), WordSelectForDesiredM (for _(M0)) = 0xFFFF (Read all 16 _(M0)words), WordSelectForDesiredM (for _(M1)) = 0xFFFF(Read all 16 _(M1)words), RE = R_(A) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per input [MSelect] and [WordSelectForDesiredM], R_(B) = R_(L), SIG_(B) = SIGout Refer to Section 15.3.1 3 A.Random None R_(A) = RL 4 C.Translate InputKeyRef = n2, DataLength (in words) = length of MWords in words as per Seq No 2 preformatted as per Section n 17.1, Data = MWords as returned from Seq No 2 preformatted as per Section 17.1, RE = R_(B), SIGE = SIG_(B) OutputKeyRef = n3, RE2 = R_(A) If ResultFlag = Pass then R_(C1) = RL2, SIG_(C) = SIGOut Refer to FIG. 15.3.1 5 A.Test KeyRef = n2, DataLength = length of MWords in words as per Seq No 2 preformatted as per Section 16.1, Data = MWords as returned from Seq No 2 parameter preformatted as per Section 16.1, RE = R_(C1), SIGE = SIG_(C) ResultFlag = Pass/Fail Random-SignM-Random-Translate-WriteFieldAuth sequence to write the new preauth value using the translating QA Device C 6 C.Random None R_(C2) = R_(L) 7 A.SignM KeyRef = n2, FieldSelect = Select bit corresponding to Pre authfield, FieldVal = new value of preauth field, ChipId = ChipId of B, R_(E) = R_(C2) If ResultFlag = Pass then R_(A1) = R_(L) SIG_(A) = SIGout Refer to Section 27.1.3.1 8 B.Random None R_(B1) = R_(L) 9 C.Translate InputKeyRef = n3, DataLength (in words) = length in words as per Seq 7 [FieldSelect] preformatted as per Section 17.1, Data = same as Seq 7 [FieldVal] preformatted as per Section 17.1, RE = R_(A1), SIGE = SIG_(A), OutputKeyRef = n2, RE2 = R_(B1) If ResultFlag = Pass then R_(C3) = R_(L2), SIG_(C) = SIGOut Refer to FIG. 15.3.1 10 B.WriteFieldsAuth KeyRef = n1, FieldNum = same as Seq 7 [FieldSelect], FieldData = same as Seq 7 [FieldVal], RE = R_(C3), SIGE = SIG_(C) ResultFlag = Pass/Fail, 31.5 Upgrading the Printer Parameters

This sequence is performed when a printer's operating parameter is upgraded.

The Parameter Upgrader QA Device stores the upgrade value which is copied to the operating parameter field of the Printer QA Device, and the count-remaining associated with upgrade value is decremented by 1 in the Parameter Upgrader QA Device.

The Parameter Upgrader QA Device output the data and signature only after completing all necessary checks for the upgrade.

31.5.1 Basic

The basic upgrade is used when the Parameter Upgrader QA Device and Printer QA Device being upgraded share a common key or a variant key i.e B.K_(n1)=A.K_(n2) or B.K_(n1)=FormKeyVariant(A.K_(n2), B.ChipID), where B is the Printer QA Device and A is the Parameter Upgrader QA Device.

Therefore, the messages and their signatures, generated by each of them can be correctly interpreted by the other.

The transfer sequence is performed using Random-Read-Random-XferField-WriteFieldsAuth. Table 320 shows the command sequence for a basic upgrade.

TABLE 320 Basic upgrade command sequence Seq No Function Parameter Random-Read-Random-XferField-WriteFieldsAuth reads M0 and M1 of the QA Device being upgraded, Parameter Upgrader QA Device produces the upgrade value for FieldNumE and Sequence data fields SEQ_1 and SEQ_2, then these values are written to the Printer QA Device. 1 A.Random None R_(A) = R_(L) 2 B.Read KeyRef = n1, SigOnly = 0, MSelect = 3 (indicates _(M0) and _(M1)), KeyIdSelect = 0x00 (no KeyIds required), WordSelectForDesiredM (for _(M0)) = 0xFFFF (Read all _(M0)words), WordSelectForDesiredM (for _(M1)) = 0xFFFF (Read all _(M1)words), RE = R_(A) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs, as per input [MSelect] and [WordSelectForDesiredM], R_(B) = RL,. SIG_(B) = SIGout Refer to Section 15.3.1 3 B.Random None R_(B1) = R_(L) 4 A.XferField KeyRef = n2, _(M0)OfExternal = First 16 words of MWords, _(M1)OfExternal = Last 16 words of MWords, ChipId = ChipId of B, FieldNumL = The field storing the upgrade value in the Parameter Upgrader QA Device. The value of this field will be copied to FieldNumE. FieldNumE = The field which will be upgraded in the Printer QA Device. RE = R_(B), R_(E2) = R_(B1), SIG_(E) = SIG_(B) If ResultFlag = Pass then FieldSelectB1 = FieldSelect − Select bits for FieldNumE and Seq data fields SEQ_1 and SEQ_2 field, FieldValB1 = FieldVal − New Value for FieldNumE (Copied from FieldNumL of the Parameter Upgrader QA Device) and sequence data fields R_(A1) = R_(L2), SIG_(A) = SIGout = Refer to Section 27.1.3.1. 5 B.WriteFieldsAuth KeyRef = n1, FieldSelect = FieldSelectB1, FieldData = FieldValB1, RE = R_(A1), SIGE = SIG_(A) ResultFlag = Pass/Fail 31.5.2 Using the Translate Function

The upgrade through the Translate function is used when the Parameter Upgrader QA Device and the Printer QA Device don't share a key between them. The translating QA Device shares a key with the Parameter Upgrader QA Device and a second key with the Printer QA Device. Therefore the messages and their signatures, generated by the Parameter Upgrader QA Device and the Printer QA Device are translated appropriately by the translating QA Device. The translating QA Device validates the Read from the Printer QA Device, and translates it for input to the XferField function. The translating QA Device will validate the output from the XferField function, and then translate it for input to WriteFieldsAuth message of the Printer QA Device.

For validating signatures using translation:

-   -   The Parameter Upgrader QA Device (A) and the translating QA         Device (C) must share a common or a variant key i.e         C.K_(n3)=A.K_(n2) or C.K_(n3)=FormKeyVariant(A.K_(n2),         C.ChipID).     -   The Printer QA Device (B) and the translating QA Device (C) must         share a common or a variant key i.e C.K_(n2)=B.K_(n1) or         B.K_(n1)=FormKeyVariant(C.K_(n2), B.ChipID).

Table 321 shows the command sequence for a basic refill using translation.

TABLE 321 An upgrade with translate command sequence Seq No Function Command Random-Read-Random-Translate-Random-XferField- Random-Translate-Random-WriteFieldsAuth reads M0 and M1 of the Printer QA Device using the translating QA Device C and then does a write of the upgrade value to FieldNumE and new sequence data to the seq data fields SEQ_1 and SEQ_2 field of the Printer QA Device using the translating QA Device C. 1 C.Random None R_(C) = R_(L) 2 B.Read KeyRef = n1, SigOnly = 0, MSelect = 0x03(indicates _(M0) and _(M1)), KeyIdSelect = 0x00 (no KeyIds required), WordSelectForDesiredM (for _(M0)) = 0xFFFF(Read all _(M0)words), WordSelectForDesiredM (for _(M1)) = 0xFFFF(Read all _(M1)words), R_(E) = R_(C) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per input [MSelect] and [WordSelectForDesiredM], R_(B) = RL, SIG_(B) = SIGout Refer to Section 15.3.1 3 A.Random None R_(A) = R_(L) 4 C.Translate InputKeyRef = n2, DataLength = MWords length in words as per Seq No 2 preformatted as per Section 17.1, Data = MWords as returned from Seq No 2 preformatted as per Section 17.1, RE = R_(B), SIGE = SIG_(B), OutputKeyRef = n3, RE2 = R_(A) If ResultFlag = Pass then R_(C1) = RL2, SIG_(C) = SIGOut Refer to Section 17.3.1 5 C.Random None R_(C2) = R_(L) 6 A.XferField KeyRef = n2, _(M0)OfExternal = First 16 words of MWords, _(M1)OfExternal = Last 16 words of MWords, ChipId = ChipId of B, FieldNumL = The field storing the upgrade value in the Parameter Upgrader QA Device. FieldNumE = The field which will be upgraded in the Printer QA Device. R_(E) = R_(C1), R_(E2) = R_(C2), SIG_(E) = SIG_(C) If ResultFlag = Pass then FieldSelectB1 = FieldSelect − Select bits for FieldNumE and sequence fields, FieldValB1 = FieldVal − New Value for FieldNumE (Copied from FieldNumL of the Parameter Upgrader QA Device) and sequence fields SEQ_1 and SEQ_2, R_(A1) = R_(L2), SIG_(A) = SIGout Refer to Section 27.1.3.1 7 B.Random None R_(B1) = R_(L) 8 C.Translate InputKeyRef = n3, DataLength = FieldValB1 length in words as per Seq No 6 preformatted as per Section 17.1, Data = FieldValB1 as returned from Seq No 6 preformatted as per Section 17.1, RE = R_(A1), SIGE = SIG_(A), OutputKeyRef = n2, RE2 = R_(B1) If ResultFlag = Pass then R_(C3) = R_(L2), SIG_(C) = SIGOut Refer to Section 17.3.1 19 B.WriteFieldsAuth KeyRef = n1, FieldSelect = FieldSelectB1, FieldVal = FieldValB1, RE = R_(C3), SIGE = SIG_(C) 10 ResultFlag = Pass/Fail 31.6 Recovering From a Failed Upgrade

This sequence is performed if the upgrade failed (for e.g Printer QA Device didn't receive the upgrade message correctly and hence didn't upgrade successfully). The Parameter Upgrader QA Device therefore needs to be rolled back to the previous value before the upgrade. In this case, the count-remaining associated with the upgrade value in the Parameter Upgrader QA Device is increased by one.

The Parameter Upgrader QA Device checks that the Printer QA Device didn't actually receive the message correctly using the StartRollBack function. The RollBackField performs further comparisons on sequence fields and FieldNumE of the Printer QA Device to values stored in the XferEntry cache. After performing all checks, the Parameter Upgrader QA Device increments the count remaining field associated with the upgrade value field by one. Refer to Section 26 and Section 28 for details.

The rollback is started using the Random-Read-Random-StartRollBack-WriteFieldsAuth and the rollback of the Parameter Upgrader QA Device is performed using Random-Read-RollBackField sequence.

Seq No Function Command Random-Read-Random-StartRollBack-WriteFieldsAuth starts the rollback and updates data for the sequence fields. 1 A.Random None R_(A) = RL 2 B.Read KeyRef = n1 , SigOnly = 0, MSelect = 0x03(indicates _(M0) and _(M1)), KeyIdSelect = 0x00 (no KeyIds required), WordSelectForDesiredM (for _(M0)) = 0xFFFF(Read all _(M0)words), WordSelectForDesiredM (for _(M1)) = 0xFFFF (Read all _(M1)words), R_(E) = R_(A) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per input [MSelect] and [WordSelectForDesiredM], R_(B) = R_(L), SIG_(B) = SIGout Refer to Section 15.3.1 3 B.Random None R_(B1) = R_(L) 4 A.StartRollBack KeyRef = n2, _(M0)OfExternal = First 16 words of MWords, _(M1)OfExternal = Last 16 words of MWords, ChipId = ChipId of B, FieldNumE = The field which was not upgraded in the Printer QA Device, FieldNumL = The upgrade value in the Parameter Upgrader QA Device which couldn't be copied to FieldNumE of the Printer QA Device, R_(E) = R_(B), R_(E2) = R_(B1), SIG_(E) = SIG_(B) If ResultFlag = Pass then FieldSelectB = FieldSelect − Select bits for sequence data fields SEQ_1 and SEQ_2, FieldValB = FieldVal − New values for SEQ_1 and SEEQ_2 fields R_(A1) = R_(L2) SIG_(A) = SIGout Refer to Section 27.1.3.1. 5 B.WriteFieldsAuth KeyRef = n1, FieldSelect = FieldSelectB, FieldData = FieldValB, RE = R_(A1), SIGE = SIG_(A) ResultFlag = Pass/Fail Random-Read-RollBackField performs a read of the QA Device being upgraded, checks its values are as per Xfer Entry cache, and then adjusts its count-remaining field. 6 A.Random None R_(A2) = RL 7 B.Read KeyRef = n1, SigOnly = 0, MSelect = 0x03 (indicates _(M0) and _(M1)), KeyIdSelect = 0x00 (no KeyIds required), WordSelectForDesiredM (for _(M0)) = 0xFFFF (Read all _(M0)words), WordSelectForDesiredM (for _(M1)) = 0xFFFF (Read all _(M1)words), R_(E) = R_(A2) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per input [MSelect] and [WordSelectForDesiredM], R_(B2) = RL, SIG_(B) = SIGout Refer to Section 15.3.1 8 A.RollBackField KeyRef = n2, _(M0)OfExternal = First 16 words of MWords, _(M1)OfExternal = Last 16 words of MWords, ChipId = ChipId of B, FieldNumE = The field which was not upgraded in the Printer QA Device, FieldNumL = The upgrade value in the Parameter Upgrader QA Device which couldn't be copied to FieldNumE of the Printer QA Device, R_(E) = R_(B2), SIG_(E) = SIG_(B) ResultFlag = Pass/Fail 31.7 Re/Filling the Consumable (INK)

This sequence is performed when an ink cartridge is first manufactured or after all the physical ink has been used, it can be filled or refilled. The re/fill protocol is used to transfer the logical ink from the Ink Refill QA Device to the Ink QA Device in the ink cartridge.

The Ink Refill QA Device stores the amount of logical ink corresponding to the physical ink in the refill station. During the refill, the required logical amount (corresponding to the physical transfer amount) is transferred from the Ink Refill QA Device to the Ink QA Device.

The Ink Refill QA Device output the transfer data only after completing all necessary checks to ensure that correct logical ink type is being transferred e.g Network_OEM1_infrared ink is not transferred to Network_OEM2_cyan ink. Refer to the XferAmount command in Section 27.1.

31.7.1 Basic Refill

The basic refill is used when the Ink Refill QA Device and the Ink QA Device share a common key or a variant key i.e B.K_(n1)=A.K_(n2) or B.K_(n1)=FormKeyVariant(A.K_(n2), B.ChipID) where B is the Ink QA Device and A is the Ink Refill QA Device. Therefore, the messages and their signatures, generated by each of them can be correctly interpreted by the other.

The Xfer Sequence is started using Random-Read-Random-StartXfer-WriteAuth and the the Xfer Amount is written to the QA Device being refilled using Random-Read-Random-XferAmount-WriteFieldsAuth sequence.

TABLE 323 the command sequence for a basic refill. Seq No Function Parameter Random-Read-Random-XferAmount-WriteFieldsAuth reads M0 and M1 of the Ink QA Device being refilled, produce updated amount for FieldNumE and sequence datat field by calling XferAmount on Ink Refill QA Device, and finally writing the updated value to Ink QA Device using WriteFieldsAuth. 1 A.Random None R_(A) = R_(L) 2 B.Read KeyRef = n1, SigOnly = 0, MSelect = 0x03(indicates _(M0) and _(M1)), KeyIdSelect = 0x00 (no KeyIds required), WordSelectForDesiredM (for _(M0)) = 0xFFFF (Read all _(M0)words), WordSelectForDesiredM (for _(M1)) = 0xFFFF(Read all _(M1)words), RE = R_(A) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per input [MSelect] and [WordSelectForDesiredM], R_(B) = RL, SIG_(B) = SIGout Refer to Section 15.3.1 3 B.Random None R_(B1) = R_(L) 4 AxferAmount KeyRef = n2, _(M0)OfExternal = First 16 words of MWords, _(M1)OfExternal = Last 16 words of MWords, ChipId = ChipId of B, FieldNumL = ink- remaining field of the Ink Refill QA Device, FieldNumE = ink- remaining field of the Ink QA Device, XferValLength = length in words of XferVal XferVal = Value to be transferred from Ink Refill QA Device to Ink QA Device being refilled, R_(E) = R_(B), R_(E2) = R_(B1), SIG_(E) = SIG_(B) If ResultFlag = Pass then FieldSelectB1 = FieldSelect − Select bits for FieldNumE and sequence data field SEQ_1 and SEQ_2, FieldValB1 = FieldVal − New Value for FieldNumE (transferred from FieldNumL of the Ink Refill QA Device) and sequence data fields SEQ_1 and SEQ_2, R_(A1) = R_(L2), SIG_(A) = SIGout Refer to Section 27.1.3.1. 5 B.WriteFieldsAuth KeyRef = n1, FieldSelect = FieldSelectB, FieldData = FieldValB, RE = R_(A1), SIGE = SIG_(A) ResultFlag = Pass/Fail 31.7.2 Using the Translate Function

The refill through the Translate function is used when the Ink Refill QA Device and the Ink QA Device don't share a key between them. The translating QA Device shares a key with the Ink Refill QA Device and a second key with the Ink QA Device. Therefore the messages and their signatures, generated by the Ink Refill QA Device and the Ink QA Device, are translated appropriately by the translating QA Device. The translating QA Device validates the Read from the Ink QA Device, and translates it for input to the XferAmount function. The translating QA Device will validate the output from the XferAmount function, and then translate it for input to WriteFieldsAuth message of the Ink QA Device.

For validating signatures using translation:

-   -   The Ink Refill QA Device (A) and the translating QA Device (C)         must share a common or a variant key i.e C.K_(n3)=A.K_(n2) or         C.K_(n3)=FormKeyVariant(A.K_(n2), C.ChipID).     -   The Ink Refill QA Device being refilled (B) and the translating         QA Device (C) must share a common or a variant key i.e         C.K_(n2)=B.K_(n1) or B.K_(n1)=FormKeyVariant(C.K_(n2),         B.ChipID).

TABLE 324 A basic refill using translation command sequence Seq No Function Command Random-Read-Random-Translate-Random-XferAmount-Random- Translate-Random-WriteFieldsAuth - reads M0 and M1 of the Ink QA Device being refilled using the translating QA Device C, produce updated amount for FieldNumE and sequence data field by calling XferAmount on Ink Refill QA Device, and finally writing the updated value to Ink QA Device using the translating QA Device. 1 C.Random None R_(C) = R_(L) 2 B.Read KeyRef = n1, SigOnly = 0, MSelect = 0x03(indicates _(M0) and _(M1)), KeyIdSelect = 0x00 (no KeyIds required), WordSelectForDesiredM (for _(M0)) = 0xFFFF (Read all _(M0)words), WordSelectForDesiredM (for _(M1)) = 0xFFFF(Read all _(M1)words), R_(E) = R_(C) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per input [MSelect] and [WordSelectForDesiredM], R_(B) = R_(L), SIG_(B) = SIGout Refer to Section 15.3.1 3 A.Random None R_(A) = R_(L) 4 C.Translate InputKeyRef = n2, DataLength = MWords length in words as per Seq No 2 preformatted as per Section 17.1, Data = MWords as returned from Seq No 2 preformatted as per Section 17.1, RE = R_(B), SIGE = SIG_(B), OutputKeyRef = n3, RE2 = R_(A) If ResultFlag = Pass then R_(C1) = R_(L2), SIG_(C) = SIGOut Refer to Section 17.3.1 5 C.Random None R_(L) = R_(C2) 6 A.XferAmount KeyRef = n2, _(M0)OfExternal = First 16 words of MWords, _(M1)OfExternal = Last 16 words of MWords, ChipId = ChipId of B, FieldNumL = ink-remaining field of the Ink Refill QA Device, FieldNumE = ink-remaining field of the Ink QA Device, XferValLength = length in words of XferVal XferVal = Value to be transferred from Ink Refill QA Device to Ink QA Device being refilled, R_(E) = R_(C1), R_(E2) = R_(C2), SIG_(E) = SIG_(C) If ResultFlag = Pass then FieldSelectB1 = FieldSelect − Select bits for FieldNumE and sequence data field SEQ_1 and SEQ_2, FieldValB1 = FieldVal − New Value or FieldNumE (transferred from FieldNumL of the Ink Refill QA Device) and sequence data fields SEQ_1 and SEQ_2, R_(A1) = R_(L2), SIG_(A) = SIGout Refer to Section 27.1.3.1 7 B.Random None R_(B1) = R_(L) 8 C.Translate InputKeyRef = n3, DataLength = FieldValB length in words as per Seq No 6 preformatted as per Section 17.1, Data = FieldValB as returned from Seq No 6 preformatted as per Section 17.1, RE = R_(A1), SIGE = SIG_(A), OutputKeyRef = n2, RE2 = R_(B1) If ResultFlag = Pass then R_(C3) = RL2, SIG_(C) = SIGOut Refer to Section 17.3.1 9 B.WriteFieldsAuth KeyRef = n1, FieldSelect = FieldSelectB, FieldData = FieldValB, RE = R_(C3), SIGE = SIG_(C) ResultFlag = Pass/Fail 31.8 Recovering From a Failed Refill

This sequence is performed if the refill failed (for e.g Ink QA Device didn't receive the refill message correctly and hence didn't refill successfully). The Ink Refill QA Device therefore needs to be rolled back to the previous value before the refill.

The Ink Refill QA Device checks that the Ink QA Device didn't actually receive the message correctly using the StartRollBack function. The RollBackAmount performs further comparisons on sequence data field and FieldNumE of the Ink QA Device, to values stored in the XferEntry cache. After performing all checks, the Ink Refill QA Device adjusts its ink field to a previous value before the transfer request was processed by it. Refer to Section 26 and Section 28 for details.

The rollback is started using the Random-Read-Random-StartRollBack-WriteFieldsAuth and the rollback of the Ink Refill QA Device is performed using Random-Read-RollBackAmount sequence.

TABLE 325 Rollback amount command sequence Seq No Function Command Random-Read-Random-StartRollBack-WriteAuth starts the rollback and updates data for the sequence data fields SEQ_1 and SEQ_2. 1 A.Random None R_(A) = RL 2 B.Read KeyRef = n1, SigOnly = 0, MSelect = 0x03(indicates _(M0) and _(M1)), KeyIdSelect = 0x00 (no KeyIds required), WordSelectForDesiredM (for _(M0)) = 0xFFFF (Read all _(M0)words), WordSelectForDesiredM (for _(M1)) = 0xFFFF(Read all _(M1)words), R_(E) = R_(A) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per input [MSelect] and [WordSelectForDesiredM], R_(B) = RL, SIG_(B) = SIGout Refer to Section 15.3.1 3 B.Random None R_(B1) = R_(L) 4 A.StartRollBack KeyRef = n2, _(M0)OfExternal = First 16 words of MWords, _(M1)OfExternal Last 16 words of MWords, ChipId = ChipId of B, FieldNumL = ink-remaining field of the Ink Refill QA Device which will be adjusted to the value before the failed refill, FieldNumE = ink-remaining field of the Ink QA Device which failed to refill, R_(E) = R_(B), R_(E2) = R_(B1) SIG_(E) = SIG_(B) If ResultFlag = Pass then FieldSelectB = FieldSelect − Select bits for sequence data fields- SEQ_1 and SEQ_2, FieldValB = FieldVal − New value for sequence data fields SEQ_1 and SEQ_2 R_(A1) = R_(L2), SIG_(A) = SIGout Refer to Section 27.1.3.1. 5 B.WriteFieldsAuth KeyRef = n1, FieldSelect = FieldSelectB in Seq No 4, FieldData = FieldValB in Seq No 4 RE = R_(A1), SIGE = SIG_(A) 10 ResultFlag = Pass/Fail Random-Read-RollBackAmount performs a read of the Ink QA Device, checks its values are as per Xfer Entry cache, and then adjusts its ink-remaining field. 11 A.Random None R_(A2). = RL 12 B.Read KeyRef = n1, SigOnly = 0, MSelect = 0x03(indicates _(M0) and _(M1)), KeyIdReq = 0 (not required), KeyIdSelect = 0x00 (no KeyIds required), WordSelectForDesiredM (for _(M0)) = 0xFFFF (Read all _(M0)words), WordSelectForDesiredM (for _(M1)) = 0xFFFF(Read all _(M1)words), RE = R_(A2) If ResultFlag = Pass then MWords = SelectedWordsOfSelectedMs as per input [MSelect] and [WordSelectForDesiredM], R_(B2) = R_(L), SIG_(B) = SIGout Refer to Section 15.3.1 13 A.RollBackAmount KeyRef = n2, _(M0)OfExternal = First 16 words of MWords, _(M1)OfExternal = Last 16 words of MWords, ChipId = ChipId of B, FieldNumL = ink- remaining field of Ink Refill QA Device which will be adjusted to the value before the failed refill, FieldNumE = ink-remaining field of Ink QA Device which failed to refill, R_(E) = R_(B2), SIG_(E) = SIG_(B) ResultFlag = Pass/Fail 31.9 Upgrading/Refilling/Filling the Upgrader

This sequence is performed when a count-remaining field in the Parameter QA Device must be updated or when the ink-remaining field in the Ink Refill QA Device requires re/filling.

In case of the Parameter QA Device, another Parameter Upgrader Refill QA Device transfers its count-remaining value to the Parameter QA Device using the transfer sequence described in Section 31.4. Also refer to Section 28.6. This means the count-remaining in the Paramater Upgrader Refill QA Device must be decremented by the same amount that Parameter Upgrader QA Device is incremented by i.e a credit transfer occurs.

In case of the Ink Refill QA Device, another Ink Refill QA Device transfers its ink-remaining value to the Ink Refill QA Device using the transfer sequence described in Section 31.4. Also refer to Section 26.4. This means the logical ink-remaining in the Ink Refill QA Device must be decremented by the same amount that QA Device being refilled is incremented by i.e a credit transfer occurs.

32 Setting Up for Field Use

This section consists of setting up the data structures in the QA Device correctly for field use. All data structures are first programmed to factory values. Some of the data structures can then be changed to application specific values at the ComCo or the OEM, while others are set to fixed values.

32.1 Instantiating the QA Chip Logical Interface

This sequence is performed when the QA Device is first created. Table 326 shows the data structure on final program load.

TABLE 326 Data structure set up during final program load Data Structure Fixed or Name Value Set to Updatable ChipId Unique Identifier for QA Device Fixed NumKey Number of keys the QA Device Fixed can hold K_(n) All K_(n) = K_(batch). The K _(batch) is Updateable if unique for a production batch^(a). previous value is known KeyId All KeyIds = KeyId of K_(batch). Updateable along with K_(n). KeyLock All KeyLock = unlocked Updateable NumVectors Number of memory vectors in the Fixed QA Device. _(M0) Set to zeros Updateable _(M0) Set to zeros Updateable M₂₊ Set to zeros Updateable P_(n) Set to ones Updateable R Set to an initial random value Updateable

Each key slot has the same K_(batch). If each key slot had a different K_(batch), and any one of the K_(batch) was compromised then the entire batch would be compromised till the K_(batch) was replaced to another key. Hence, each key slot having a different K_(batch) doesn't have any security advantages but requires more keys to be managed.

32.2 Setting Up Application Specific Data

The section defines the sequences for configuring the data structures in the QA Device to application specific data.

32.2.1 Replacing Keys

The QA Devices are programmed with production batch keys at final program load. The COMCO keys replace the production batch keys before the QA Devices are shipped to the ComCo. The ComCo replaces the COMCO keys to COMCO_OEM when shipping QA Devices to its OEMs. The OEM replaces the COMCO_OEM to COMCO_OEM app as the QA Devices are placed in ink cartridges or printers.

The replacement occurs without the ComCo or the OEM knowing the actual value of the key. The actual value of the keys is only to known to QACo. The ComCo or the OEM is able to perform these replacements because the QACo provides them with a key programming QA Device with keys appropriately set which can generate the necessary messages and signatures to replace the old key with the new key.

Table 327 shows the command sequence for ReplaceKey. The GetProgramKey gets the new encrypted key from the key programming QA Device, and the encrypted new key is passed into the QA Device whose key is being replaced through the ReplaceKey function. Depending on the OldKeyRef and NewKeyRef objects a common encrypted key or a variant encrypted key can be produced for the ReplaceKey function

TABLE 327 ReplaceKey command sequence Seq No Function Command 1 B.Random None R_(B) = R_(L) 2 A.GetProgramKey OldKeyRef = Key Num of the old key. This key must be changed to the NewKeyRef in the QA Device whose key s being replaced. ChipId = Chip identifier of the QA Device whose key is being replaced. RE = R_(B) KeyLock = Set depending on whether the new key is the final key for the key slot or it will be replaced further. NewKeyRef = Key Num of the new key. This key will change the OldKeyRef in the QA Device whose key is being replaced. If ResultFlag = Pass then R_(A) = RL, KeyId_(new) = KeyIdOfNewKey EncryptedNewKey = EncryptedKey SIGA = SIGout Refer to Section 22.2.1. 3 B.ReplaceKey KeyNumToBeReplaced = Old key number, the old key could be a common key or a variant key, KeyId = KeyId_(new), EncryptedKey = EncryptedNewKey, RE = RA, SIGE = SIGA ResultFlag = Pass/Fail 32.2.2 Setting up ReadOnly Data

This sets the permanent functional parameters of the application where the QA Device has been placed. These parameters remain unchanged for the lifetime of the QA Device. In case of the ink cartridge such parameters are colour and viscosity of the ink. These values are written to M₂₊ memory vectors using the WriteM1+ function, and its permissions are set to ReadOnly by SetPerm function. These values are typically set at the OEM.

TABLE 328 ReadOnly data setup command sequence Seq No Function Command 1 B.WriteM1+ VectNum = 2 or 3, WordSelect = the selected words to be written, MVal = words corresponding to word select starting from LSW ResultFlag = Pass/Fail 2 B.SetPerm (VectNum = same as Seq No 1parameter [VectNum], PermVal = same as Seq No 1 parameter [WordSelect]) If ResultFlag = Pass then CurrPerm = NewPerm Current permission value after applying PermVal

In case of the SBR4320, the values written to M₂₊ memory vectors is write-once only i.e they are set to ReadOnly as soon as they are written to once, therefore the command sequence consists only of Seq No 1 in Table 329.

32.2.3 Defining Fields in _(M0)

The QACo must determine the field definitions for M0 depending on the application of the QA Device. These field definitions will consist of the following:

-   -   Number of fields and the size of each field.     -   The Type attribute of each field.     -   The access permission for each field.

Following fields have been presently defined in an ink QA Device:

-   -   ink-remaining field. See Section 26 for details.     -   Preauthorisation field. See Section 31.4.3 for details.     -   Sequence data fields SEQ_(—)1 and SEQ_(—)2. See Section 26 for         details.

Following fields have been presently defined in a printer QA Device:

-   -   Operating parameter field. See Section 28 for details.     -   Sequence data fields SEQ_(—)1 and SEQ_(—)2. See Section 26 for         details.

After the field definitions are determined, they are formatted as per Section 8.1.1.4. These formatted values are then written to _(M1) using a WriteM1+ function.

TABLE 329 Defining M0 fields command sequence Sequence No Function Command 1 B.WriteM1+ VectNum = 1, WordSelect = The selected words corresponding to the attribute field/fields of _(M0), MVal = words corresponding to word select starting from LSW) ResultFlag = Pass/Fail 32.2.4 Writing Values to Fields in _(M0)

The writing of _(M0) fields for an Ink QA Device will typically occur when the ink cartridge is filled with physical ink for the first time, and the equivalent logical ink is written to the Ink QA Device. Refer to Section 31.7 for details.

The writing of _(M0) fields for a Printer QA Device will typically occur when the printer parameters are written for the first time. The procedure for writing of a printer parameter for the first time or upgrading a printer parameters is exactly the same. Refer to Section 31.5 for details.

Before any value is written to a field, the key slot containing the key which has authenticated ReadWrite access to the field must be locked.

Both Ink QA Device and Printer QA Device has a sequence data fields SEQ_(—)1 and SEQ_(—)2 as described in Section 27. These two fields must be initialised to 1xFFFFFFFF, refer to Section 27 for details.

The Ink QA Device/Printer QA Device and the trusted QA Device writing to it, share the sequence key or a variant sequence key between them i.e B.K_(n1)=A.K_(n2) or B.K_(n1)=FormKeyVariant(A.K_(n2), B.ChipID), where B is the Ink QA Device/Printer QA Device and A is the trusted QA Device. The command sequence used is described in Table 330.

TABLE 330 Command sequence for writing sequence data fields to the QA Devices. Sequence No Function Parameters 1 B.Random R_(B) = R_(L) 2 A.SignM KeyRef = n2, FieldSelect = Select bit correponding to SEQ_1 and SEQ-2 FieldVal = both fields set 0xFFFFFFFF. Refer to Section 31.4.3.3 ChipId = ChipId of B, R_(E) = R_(B) If ResultFlag = Pass then R_(A) = R_(L) SIG_(A) = SIGout Refer to Section 27.1.3.1 3 B.WriteFieldsAuth KeyRef = n1, FieldSelect = same as Seq 2[FieldSelect], FieldVal = same as Seq 2[FieldVal], RE = R_(A), SIGE = SIG_(A) ResultFlag = Pass/Fail 32.3 Setting Up the Upgrading QA Device

The upgrading QA Device must be set up either as an Ink Refill QA Device or as a Parameter Upgrader QA Device.

Each upgrading QA Device must go through the following set up:

-   -   The upgrading QA Device must be set to factory defaults. Refer         to Section 32.1. At the end of this process the upgrading QA         Device is either an Ink Refill QA Device or a Parameter Upgrader         QA Device with production batch keys and M0 fields set to         deafult.     -   The upgrading QA Device must be programmed with the appropriate         keys and upgrade data before it can start upgrading other QA         Devices. Following must be performed on each upgrade QA Device:

-   a. The upgrading QA Device must be programmed with the appropriate     keys required to upgrade other QA Devices and to upgrade itself when     necessary.

-   b. The M0 fields must be correctly defined and set in M1. For a Ink     Refill QA Device the ink-remaining field must be defined and set.     For a printer upgrade QA Device the upgrade value field and the     count-remaining field must be defined and set. All upgrade QA     Devices must also have a sequence datat fields SEQ_(—)1 and SEQ_(—)2     which are used to upgrade the upgrading QA Device itself.

-   c. Finally, M0 fields defined in b must be written with appropriate     values so that the upgrade QA Device can perform upgrades.

An Ink Refill QA Device will typically store the logical ink equivalent to the physical ink in a refill station, hence the Ink Refill QA Device's ink-remaining field must be written with the equivalent logical ink amount.

For a Parameter Upgrader QA Device the upgrade value field and the count-remaining field must be written. The upgrade value depends on the type of upgrade the Parameter Upgrader QA Device can perform i.e one Parameter Upgrader QA Device can upgrade to 10 ppm (pages per minute) while another Parameter Upgrader QA Device can upgrade to 5 ppm. The count-remaining is the number of times the Parameter Upgrader QA Device is permitted to write the associated upgrade value to other QA Devices. The count-remaining field must be written to a positive non-zero value for the Parameter Upgrader QA Device to perform successful upgrades. Refer to Section 32.3.1 and Section 32.3.2 for details.

32.3.1 Setting Up the Ink Refill QA Device

32.3.1.1 Setting Up the Keys

The Ink Refill QA DeviceQA Device could be transferring ink between peers or transferring ink down the heirachy, accordingly the peer to peer Ink Refill QA Device has two keys (fill/refill key and sequence key) as described in Section 27, and a Ink Refill QA Device transferring down the heirachy has three keys (fill/refill key, transfer key and sequence key). These keys must be programmed into the Ink Refill QA Device using the sequence described in Section 32.2.1.

The Key Programming QA Device must be programmed with the appropriate production batch keys, and the fill/refill, transfer key and sequence key

The GetProgramKey function is called on the Key Programming QA Device with OldKeyRef (OldKeyRef—refer to Section 32.2.1) pointing to a production batch key, and the NewKeyRef (NewKeyRef—refer to Section 32.2.1) pointing to either a fill/refill key or a transfer key or a sequence key. The outputs from the GetProgramKey (signature and encrypted New Key) is passed in to ReplaceKey function of the Ink Refill QA Device.

The GetProgramKey function must be called (on the Key Programming QA Device) for replacing each of the production batch keys in the Ink Refill QA Device. The output of the GetProgramKey will be passed in to the ReplaceKey function called on the Ink Refill QA Device. The successful processing of the ReplaceKey function will replace an old key(production keys) to a corresponding new key (either a fill/refill key or a transfer key or a sequence key).

32.3.1.2 Setting Up the M0 Field Information in _(M1)

The ink-remaining field and the sequence data fields SEQ_(—)1 and SEQ_(—)2 must be defined and set in the Ink Refill QA Device using the sequence described in Section 32.2.3.

32.3.1.3 Transferring Ink Amounts

Finally, the logical ink amounts are transferred to the ink-remaining field using the sequence described in Section 31.7.

The QACo will transfer to the ComCo Ink Refill QA Device at the top of the heirachy using the command sequence in Table 331.

For a successful transfer from QACo to ComCo, ComCo and QACo must share a common key or a variant key be i.e ComCo.K_(n1)=QACo.K_(n2) or ComCo.K_(n1)=FormKeyVariant(QACo.K_(n2), ,ComCo.ChipID)K_(n1) is the fill/refill key for the ComCo refill QA Device.

TABLE 331 Command sequence for writing ink-remaining amounts to the highest QA Device in the heirachy. Sequence No Function Parameters 1 B.Random R_(B) = RL 2 A.SignM KeyRef = n2, FieldSelect = Select bit correponding to the ink-remaining field, FieldVal = Ink amount to be transferred, Refer to Section 31.4.3.3 ChipId = ChipId of B, R_(E) = R_(B) If ResultFlag = Pass then R_(A) = R_(L) SIG_(A) = SIGout Refer to Section 27.1.3.1 3 B.WriteFieldsAuth KeyRef = n1, FieldSelect = same as Seq 2[FieldSelect], FieldVal = same as Seq 2[FieldVal], RE = R_(A), SIGE = SIG_(A) ResultFlag = Pass/Fail 32.3.1.4 Setting Up Sequence Data Fields

The Ink Refill QA Device has sequence data fields SEQ_(—)1 and SEQ_(—)2 (as described in Section 27) because its ink-remaining fields can be refilled as well. These two fields must be initialised to 0xFFFFFFFF, refer to Section 27 for details.

The Ink Refill QA Device and the trusted QA Device writing to it, share the sequence key or a variant sequence key between them i.e B.K_(n1)=A.K_(n2) or B.K_(n1)=FormKeyVariant(A.K_(n2), B.ChipID), where B is the Ink Refill QA Device and A is the trusted QA Device. The command sequence used is described in Table 331.

32.3.2 Setting Up the Parameter Upgrader QA Device

32.3.2.1 Setting Up the Keys

The Parameter Upgrader QA Device could be transferring upgrades between peers or transferring upgrades down the heirachy, accordingly the peer to peer Parameter Upgrader QA Device has three keys (write-parameter key, fill/refill key and sequence key) as described in Section 28.6 and Section 26, and a Parameter Upgrader QA Device transferring down the heirachy has four keys (write-parameter key, fill/refill key, transfer key and sequence Key). These keys must be programmed into the Parameter Upgrader QA Device using the sequence described in Section 32.2.1.

The Key Programming QA Device must be programmed with the appropriate production batch keys, and write-parameter key, fill/refill key, transfer key and sequence key

The GetProgramKey function is called on the Key Programming QA Device with OldKeyRef (OldKeyRef—refer to Section 32.2.1) pointing to a production batch key, and the NewKeyRef (NewKeyRef—refer to Section 32.2.1) pointing to either a write-parameter key, or a fill/refill key, or a transfer key, or a sequence key. The outputs from the GetProgramKey (signature and encrypted New Key) is passed in to ReplaceKey function of the Parameter Upgrader QA Device.

32.3.2.2 Setting Up the M0 Field in _(M1)

The upgrade value field and the count-remaining field must be defined and set in the upgrade QA Device using the sequence described in Section 32.2.3.

32.3.2.3 Writing Upgrade Value to the Upgrade Field

The upgrade value is written to upgrade field using the write-parameter key. The upgrade QA Device and the trusted QA Device writing to it, share the write-parameter key or a variant write-parameter key between them i.e B.K_(n1)=A.K_(n2) or B.K_(n1)=FormKeyVariant(A.K_(n2), B.ChipID), where B is the upgrade QA Device and A is the trusted QA Device. The command sequence used is described in Table 331.

32.3.2.4 Transferring Count-Remaining Amounts

Finally, the logical count-remaining amounts are transferred to the count-remaining field using the sequence described in Section 31.7.

The QACo will also transfer to the ComCo's upgrade QA Device using the command sequence in Table 331.

For a successful transfer from QACo to ComCo, ComCo and QACo must share a common key or a variant key be i.e ComCo.K_(n1)=QACo.K_(n2) or ComCo.K_(n1)=FormKeyVariant(QACo.K_(n2),ComCo.ChipID). K_(n1) is the fill/refill key for the ComCo upgrade QA Device.

32.3.2.5 Setting Up Sequence Data Fields

The Parameter Upgrader QA Device has sequence data fields SEQ_(—)1 and SEQ_(—)2 (as described in Section 27) because its count-remaining fields can be refilled as well. These two fields must be initialised to 0xFFFFFFFF, refer to Section 27 for details.

The Parameter Upgrader QA Device and the trusted QA Device writing to it, share the sequence key or a variant sequence key between them i.e B.K_(n1)=A.K_(n2) or B.K_(n1)=FormKeyVariant(A.K_(n2), B.ChipID), where B is the Parameter Upgrader QA Device and A is the trusted QA Device. The command sequence used is described in Table 331.

32.4 Setting Up the Key Programmer

The key programming QA Device is set up to replace keys in other QA Devices.

Each key programming QA Device must go through the following set up:

-   -   The key programming QA Device must be instantiated to factory         defaults. Refer to Section 32.1. At the end of instantiation the         key programming QA Device has production batch keys and no key         replacement data.     -   The key programming QA Device must be programmed with the         appropriate keys and key replacement map before it can start to         replace keys in other QA Devices.         32.4.1 Setting Up the Keys

The key programming QA Device must be programmed with the key replacement map key. The key replacement map key is described in details in Section 24.

The key programming QA Device must programmed with the old and new keys for the QA Devices it is going to perform key replacement on.

Each of the keys is set in the key programming QA Device using the sequence described in Section 32.2.1.

32.4.2 Setting Up Key Replacement Map Field Information

First the key replacement map field information is worked out as per Section 24.1. This field information is set in M1 as per the sequence described Section 32.2.3.

32.4.3 Setting Up Key Replacement Map

Finally, the key replacement map field must be written with the valid mapping using the key replacement map key. The key programming QA Device and the trusted QA Device writing to it must share the key replacement map key or a variant of the key replacement map key between them.

For a successful write of the key replacement map B.K_(n1)=A.K_(n2) or B.K_(n1)=FormKeyVariant(A.K_(n2), B.ChipID), where B is the key replacement QA Device and A is the trusted QA Device. The command sequence used is described in Table 331.

Appendix A: Field Types

Table 332 lists the field types that are specifically required by the QA Chip Logical Interface and therefore apply across all applications. Additional field types are application specific, and are defined in the relevant application documentation.

TABLE 332 Predefined Field Types Value Type Description 0x0000 0 Non-initialised (default value after final program load) 0x0001 TYPE_PREAUTH Defines a preauth field in an Ink QA Device 0x0002 TYPE_COUNT_REMAINING Defines a countRemaining field in an Parameter Upgrader QA Device 0x0003 TYPE_SEQ_1 Defines a sequence data field SEQ_1 in an Ink QA Device or in a Printer QA Device or in an upgrader QA Device 0x0004 TYPE_SEQ_2 Defines a sequence data fields SEQ_2 in an Ink QA Device or in a Printer QA Device or in an upgrader QA Device 0x0005 TYPE_KEY_MAP Defines a key replacement map in a Key Programmer QA Device 0x0006 reserved reserved for and future use above

Appendix B: Key and field definition for different QA Devices

B.1 Parameter Upgrader QA Device

B.1.1 Peer to Peer QA Device

TABLE 333 Key definitions for a peer to peer Parameter Upgrader QA Device Key Name Purpose Fill/refill This key has is used for upgrading Key count-remaining values when the upgrade QA Device is upgraded by another upgrade QA Device and is also used to decrement the count- remaining when upgrading other QA Devices. Sequence This key is used to initialise Key sequence data fields SEQ_1 and SEQ_2 to 0xFFFFFFF. Write This key is used to write the Parameter upgrade value to the Parameter Key Upgrader QA Device.

TABLE 334 Field definitions for a peer to peer Parameter Upgrader QA Device Field Attrinutes Field A^(a) NA^(b) EndPos Name Purpose Type KeyNum RW RW KPerms^(c) (Size) CountRemaining The field stores TYPE_COUNT_REMAINING SN^(f) fill/ 1 0 KPerms[KN^(e)] = 1 Depends on the number of refill key Rest are 0 the maximum times the number of Parameter upgrades Upgrader QA that can Device is be stored. permitted to upgrade a printer QA Device. UpgradeValue This stores the Must define the SN^(f) write- 1 0 KPerms[KN^(e)] = 0 Set as per value that is type of the parameter Rest are 0 upgrade copied from the upgrade value i.e key as well value. Parameter TYPE_PRINT_SPEED^(d) Upgrader QA Device to the field being upgraded on the printer QA Device during the upgrade SEQ_1 This field holds TYPE_SEQ_1 SN^(f) 1 0 KPerms[KN^(e)] = 0 Typically the data for sequence KPerms[fill/ 32 bit. sequence data key refill^(g)] = 1 field SEQ_1 Rest are 0 when the as well. Parameter Upgrader QA Device is being upgraded by another Parameter Upgrader Refill QA Device. SEQ_2 This field holds TYPE_SEQ_2 SN^(f) sequence 1 0 KPerms[KN^(e)] = 0 Typically the data for key KPerms[fill/ 32 bit. sequence data refill^(g)] = 1 fieldsSEQ_2 Rest are 0 when the as well. Parameter Upgrader QA Device is being upgraded by another Parameter Upgrader Refill QA Device. ^(a)Authenticated ReadWrite permission ^(b)Non-authenticated ReadWrite permission ^(c)KeyPerms ^(d)This is a sample type only ^(e)KeyNum ^(f)Key Slot Number ^(g)Fill/Refill key has authenticated decrement-only permission to the sequence data fields B.1.2 Heirarchical Transfer QA Device Key Definitions

TABLE 335 Key definitions for a Parameter Upgrader QA Device (transferring down the heirachy) Key Name Purpose Transfer This key is used to decrement Key the count-remaining when upgrading other QA Devices. Fill/refill This key has is used for Key upgrading count-remaining values when the Parameter Upgrader QA Device is upgraded by another Parameter Upgrader QA Device Refill QA Device. Sequence This key is used to Key initialise sequence data fields SEQ_1 and SEQ_2 to 0xFFFFFFF. Write This key is used to write Parameter the upgrade value to the Key Parameter Upgrader QA Device. Field Definitions

TABLE 336 Field definitions for Parameter Upgrader QA Device transferring down the hierachy Field Attrinutes Field A^(a) NA^(b) EndPos Name Purpose Type KeyNum RW RW KPerms^(c) (Size) Count The field stores the TYPE_COUNT_REMAINING SN^(f) fill/ 1 0 KPerms[KN^(e)] = 0 Depends on the Remaining number of times refill key KPerms[Transfer maximum number the Parameter Key] = 1 of upgrades Upgrader QA Rest are 0 that can be Device is permitted stored. to upgrade a printer QA Device. Upgrade This stores the Must define the type of SN^(f) write- 1 0 KeyPerms[KN^(e)] = 0 Set as per Value value that is the upgrade value parameter Rest are 0 upgrade value. copied from the i.e key Parameter TYPE_PRINT_SPEED^(d) Upgrader QA Device to the field being upgraded on the printer QA Device during the upgrade SEQ_1 This field holds TYPE_SEQ_1 SN^(f) sequence 1 0 KPerms[KN^(e)] = 0 Typically the data for key KPerms[fill/ 32 bit. sequence data refill^(g)] = 1 fields SEQ_1 Rest are 0 as when the well. Parameter Upgrader QA Device is being upgraded by another Parameter Upgrader Refill QA Device. SEQ_2 This field holds TYPE_SEQ_2 SN^(f) sequence 1 0 KPerms[KN^(e)] = 0 Typically the data for key KPerms[fill/ 32 bit. sequence data refill^(g)] = 1 fields SEQ_2 Rest are 0 as when the well. Parameter Upgrader QA Device is being upgraded by another Parameter Upgrader Refill QA Device. ^(a)Authenticated ReadWrite permission ^(b)Non-authenticated ReadWrite permission ^(c)KeyPerms ^(d)This is a sample type only ^(e)KeyNum ^(f)Key Slot Number ^(g)Fill/Refill key has authenticated decrement-only permission to the sequence data fields B.2 Ink Refill QA Device B.2.1 Peer to Peer QA Device Key Definitions

TABLE 337 Key definitions for a peer to peer Ink Refill QA Device Key Name Purpose Fill/refill This key has is used for filling/ Key refilling ink-remaining values when the Ink Refill QA Device is upgraded by another Ink Refill QA Device and is also used to decrement from the ink-remaining when transferring ink to other QA Devices (typically Ink QA Device). Sequence This key is used to initialise Key sequence data fields SEQ_1 and SEQ_2 to 0xFFFFFFF. Field Definitions

TABLE 338 Field definitions for a peer to peer Ink Refill QA Device Field Attrinutes Field Key A^(a) NA^(b) EndPos Name Purpose Type Num RW RW KeyPerms^(c) (Size) InkRemaining The field Must define the SN^(f) fill/ 1 1 KeyPerms[KN^(e)] = 1 Depends on the stores the type of Ink e.g refill key Rest are 0 maximum amount amount of TYPE_HIGHQUALITY_(—) of ink that logical ink- BLACK_INK^(d) can be stored remaining in and the storage the ink refill resolution i.e QA Device. in pico litres or in micro litres. SEQ_1 This field TYPE_SEQ_1 SN^(f) sequence 1 0 KPerms[KN^(e)] = 0 Typically holds the key KPerms[fill/ 32 bit. data for refill^(g)] = 1 sequence data Rest are 0 as field SEQ_1 well. when the Ink Refill QA Device is being filled/ refilled by another Ink Refill QA Device. SEQ_2 This field TYPE_SEQ_2 SN^(f) sequence 1 0 KPerms[KN^(e)] = 0 Typically holds the data key KPerms[fill/ 32 bit. for sequence refill^(g)] = 1 data field Rest are 0 as SEQ_2 when the well. Ink Refill QA Device is being filled/refilled by another Ink Refill QA Device. ^(a)Authenticated ReadWrite permission ^(b)Non-authenticated ReadWrite permission ^(c)Decrement-Only For Keys ^(d)This is a sample type only ^(e)KeyNum ^(f)Key Slot Number ^(g)Fill/Refill key has authenticated decrement-only permission to the sequence data fields B.2.2 Heirarchical Transfer QA Device Key Definitions

TABLE 339 Key definitions for a ink refill QA Device (transferring down the heirachy) Key Name Purpose Transfer This key is used to decrement from the Key ink-remaining when transferring ink to other QA Devices. Fill/refill This key has is used for filling/ Key refilling ink-remaining values when the Ink Refill QA Device is upgraded by another Ink Refill QA Device. Sequence This key is used to initialise sequence Key data fields SEQ_1 and SEQ_2 to 0xFFFFFFF. Field Definitions

TABLE 340 Field definitions for a Ink Refill QA Device (transferring down the heirachy) Field Attrinutes Field A^(a) NA^(b) EndPos Name Purpose Type KeyNum RW RW KeyPerms^(c) (Size) InkRemaining The field Must define SN^(f) fill/ 1 0 KPerms[KN^(e)] = 0 Depends on the stores the the type of refill key KPerms[Transfer maximum amount amount of Ink e.g- Key] = 1 of ink that logical ink- TYPE_HIGHQUALITY_(—) Rest are 0 can be stored remaining in BLACK_INK^(d) and the storage the Ink resolution i.e Refill QA in pico litres Device. or in micro litres. SEQ_1 This field TYPE_SEQ_1 SN^(f) sequence 1 0 KPerms[KN^(e)] = 0 Typically holds the key KPerms[fill/ 32 bit. data for refill^(g)] = 1 sequence Rest are 0. data field SEQ_1 when the Ink Refill QA Device is being filled/ refilled by another Ink Refill QA Device. SEQ_2 This field TYPE_SEQ_2 SN^(f) sequence 1 0 KPerms[KN^(e)] = 0 Typically holds the key KPerms[fill/ 32 bit. data for refill^(g)] = 1 sequence Rest are 0. data field SEQ_2 when the Ink Refill QA Device is being filled/refilled by another Ink Refill QA Device. ^(a)Authenticated ReadWrite permission ^(b)Non-authenticated ReadWrite permission ^(c)KeyPerms ^(d)This is a sample type only ^(e)KeyNum ^(f)Key Slot Number ^(g)Fill/Refill key has authenticated decrement-only permission to the sequence data fields B.3 Key Programming QA Device B.3.1 Key Definitions

TABLE 341 Key definitions for a Key Programming QA Device Key Name Purpose Key This key is used to write the key replacement map. replacement map Key Old Keys These are the old keys of the QA Device whose keys will be replaced by the Key Programming QA Device. New Keys These are the new keys of the QA Device whose old keys will be replaced by the Key Programming QA Device. B.3.2 Field Definitions

TABLE 342 Field definitions for a key replacement QA Device Field Attrinutes Field A^(a) NA^(b) EndPos Name Purpose Type KeyNum RW RW KPerms^(c) (Size) Key This defines TYPE_KEY_MAP Key Replace- 1 0 KPerms[KN^(d)] = 0 2 words replacement the mapping ment Map key Rest are 0 (64 bits) map between the old key and the new key for the QA Device whose old key will be replaced by the new key. ^(a)Authenticated ReadWrite permission ^(b)Non-authenticated ReadWrite permission ^(c)KeyPerms ^(d)KeyNum B.4 Ink QA Device B.4.1 Key Definitions

TABLE 343 Key definitions for a Ink QA Device Key Name Purpose Fill/refill key This key is used for fill/refilling ink-remaining amount in the ink QA Device. Ink usage Key This key is verifying the data read from the ink QA Device and for writing preauth data. Sequence Key This key is used to initialise sequence data fields SEQ_1 and SEQ_2 to 0xFFFFFFF. B.4.2 Field Definitions

TABLE 344 Field definitions for a Ink QA Device Field Attrinutes Field Key A^(a) NA^(b) EndPos Name Purpose Type Num RW RW KPerms^(c) (Size) Ink The amount of Must define SN^(f) fill/ 1 1 KPerms[KN^(e)] = 1 Depends on the Remaining logical ink- the type of refill key Rest are 0 maximum amount remaining in Ink i.e of ink that the ink QA TYPE_HQ_(—) can be stored Device. More BLACK_INK^(d) and the storage than one ink- resolution i.e remaining in pico litres field may or in micro be present litres. depending on the number of physical inks stored in the ink cartridge. Preauth This field TYPE_PREAUTH SN^(f) ink 0 1 KPerms[KN^(e)] = 0 Depends on defines the usage key Rest are 0 preauth amount. preauth value. Typically 32 bits, may be 64 bits to accomodate larger preauth amounts. SEQ_1 This field TYPE_SEQ_1 SN^(f) 1 0 KPerms[KN^(e)] = 0 Typically holds the sequence KPerms[fill/ 32 bit. data for key refill^(g)] = 1 sequence Rest are 0. data field SEQ_1 when the Ink QA Device is being filled/ refilled by a Ink Refill QA Device. SEQ_2 This field TYPE_SEQ_2 SN^(f) 1 0 KPerms[KN^(e)] = 0 Typically holds the sequence KPerms[fill/ 32 bit. data for key refill^(g)] = 1 sequence Rest are 0. data field SEQ_2 when the Ink QA Device is being filled/ refilled by another Ink Refill QA Device. ^(a)Authenticated ReadWrite permission ^(b)Non-authenticated ReadWrite permission ^(c)KeyPerms ^(d)This is a sample type only ^(e)KeyNum ^(f)Key Slot Number ^(g)Fill/Refill key has authenticated decrement-only permission to the sequence data fields B.5 Printer QA Device B.5.1 Key Definition

TABLE 345 Key definitions for a Printer QA Device Key Name Purpose Upgrade key This key is used for writing/upgrading the (fill/refill key) functional parameter. Ink usage Key This key is verifying the data read from the Ink QA Device. Sequence Key This key is used to initialise sequence data fields SEQ_1 and SEQ_2 to 0xFFFFFFF. PECID/SOPECID This key is used to verify the data read Key from the printer QA Device. This key is unique to each printer. Also used to translate data from the ink QA Device to the trusted printer system QA Device. B.5.2 Field Definition

TABLE 346 Field definitions for a Printer QA Device Field Attrinutes A^(a) NA^(b) EndPos Field Name Purpose Type Key Num RW RW KPerms^(c) (Size) Functional The field Must define SN^(f) fill/ 1 0 KPerms[KN^(e)] = 0 Set as per parameter stores an the type of refill key Rest are 0 functional upgradeable print speed parameter. functional i.e parameter. TYPE_PRINT_SPEED^(d) More than one functional parameter can be stored in the printer QA Device. SEQ_1 This field TYPE_SEQ_1 SN^(f) 1 0 KPerms[KN^(e)] = 0 Typically holds the sequence KPerms[fill/refill^(g)] = 1 32 bit. data for key Rest are 0. sequence data field SEQ_1 when the Printer QA Device is being filled/ refilled by a Parameter Upgrader QA Device. SEQ_2 This field TYPE_SEQ_2 SN^(f) 1 0 KPerms[KN^(e)] = 0 Typically holds the data sequence KPerms[fill/refill^(g)] = 1 32 bit. for sequence key Rest are 0. data field SEQ_2 when the Printer QA Device is being filled/ refilled by another Parameter Upgrader QA Device. ^(a)Authenticated ReadWrite permission ^(b)Non-authenticated ReadWrite permission ^(c)KeyPerms ^(d)This is a sample type only ^(e)KeyNum ^(f)Key Slot Number ^(g)Fill/Refill key has authenticated decrement-only permission to the sequence data fields B.6 Trusted Printer System QA Device B.6.1 Key Definition

TABLE 347 Key Name Purpose PECID/SOPECID This key is used to verify the data read from Key the printer QA Device. This key is unique to each printer. This key is also used for verifying translated data from the ink QA Device. Introduction 1 Background

This document describes a QA Chip that can be used to hold contains authentication keys together with circuitry specially designed to prevent copying. The chip is manufactured using a standard Flash memory manufacturing process, and is low cost enough to be included in consumables such as ink and toner cartridges. The implementation is approximately 1 mm² in a 0.25 micron flash process, and has an expected die manufacturing cost of approximately 10 cents in 2003.

Once programmed, the QA Chips as described here are compliant with the NSA export guidelines since they do not constitute a strong encryption device. They can therefore be practically manufactured in the USA (and exported) or anywhere else in the world.

Note that although the QA Chip is designed for use in authentication systems, it is microcoded, and can therefore be programmed for a variety of applications.

2 Nomenclature

The following symbolic nomenclature is used throughout this document:

TABLE 348 Summary of symbolic nomenclature Symbol Description F[X] Function F, taking a single parameter X F[X, Y] Function F, taking two parameters, X and Y X | Y X concatenated with Y X

Y Bitwise X AND Y X

Y Bitwise X OR Y (inclusive-OR) X ⊕ Y Bitwise X XOR Y (exclusive-OR)

X Bitwise NOT X (complement) X

Y X is assigned the value Y X

{Y, Z} The domain of assignment inputs to X is Y and Z X = Y X is equal to Y X ≠ Y X is not equal to Y

X Decrement X by 1 (floor 0)

X Increment X by 1 (modulo register length) Erase X Erase Flash memory register X SetBits[X, Y] Set the bits of the Flash memory register X based on Y Z

ShiftRight[X, Y] Shift register X right one bit position, taking input bit from Y and placing the output bit in Z 3 Pseudocode 3.1 Asynchronous

The following pseudocode:

-   -   var=expression         -   means the var signal or output is equal to the evaluation of             the expression.             3.2 Synchronous

The following pseudocode:

-   -   var←expression         -   means the var register is assigned the result of evaluating             the expression during this cycle.             3.3 Expression

Expressions are defined using the nomenclature in Table 348 above. Therefore:

-   -   var=(a=b)         -   is interpreted as the var signal is 1 if a is equal to b,             and 0 otherwise.             4 Diagrams

Black lines are used to denote data, while red lines are used to denote 1-bit control-signal lines.

Logical Interface

5 Introduction

The QA Chip has a physical and a logical external interface. The physical interface defines how the QA Chip can be connected to a physical System, while the logical interface determines how that System can communicate with the QA Chip. This section deals with the logical interface.

5.1 Operating Modes

The QA Chip has four operating modes—Idle Mode, Program Mode, Trim Mode and Active Mode.

-   -   Active Mode is entered on power-on Reset when the fuse has been         blown, and whenever a specific authentication command arrives         from the System. Program code is only executed in Active Mode.         When the reset program code has finished, or the results of the         command have been returned to the System, the chip enters Idle         Mode to wait for the next instruction.     -   Idle Mode is used to allow the chip to wait for the next         instruction from the System.     -   Trim Mode is used to determine the clock speed of the chip and         to trim the frequency during the initial programming stage of         the chip (when Flash memory is garbage). The clock frequency         must be trimmed via Trim Mode before Program Mode is used to         store the program code.     -   Program Mode is used to load up the operating program code, and         is required because the operating program code is stored in         Flash memory instead of ROM (for security reasons).

Apart from while the QA Chip is executing Reset program code, it is always possible to interrupt the QA Chip and change from one mode to another.

5.1.1 Active Mode

Active Mode is entered in any of the following three situations:

-   -   power-on Reset when the fuse has been blown     -   receiving a command consisting of a global id write byte (0x00)         followed by the ActiveMode command byte (0x06)     -   receiving a command consisting of a local id byte write followed         by some number of bytes representing opcode and data.

In all cases, Active Mode causes execution of program code previously stored in the flash memory via Program Mode.

If Active Mode is entered by power-on Reset or the global id mechanism, the QA Chip executes specific reset startup code, typically setting up the local id and other IO specific data. The reset startup code cannot be interrupted except by a power-down condition. The power-on reset startup mechanism cannot be used before the fuse has been blown since the QA Chip cannot tell whether the flash memory is valid or not. In this case the globalid mechanism must be used instead.

If Active Mode is entered by the local id mechanism, the QA Chip executes specific code depending on the following bytes, which function as opcode plus data. The interpretation of the following bytes depends on whatever software happens to be stored in the QA Chip.

5.1.2 Idle Mode

The QA Chip starts up in Idle Mode when the fuse has not yet been blown, and returns to Idle Mode after the completion of another mode. When the QA Chip is in Idle Mode, it waits for a command from the master by watching the low speed serial line for an id that matches either the global id (0x00), or the chip's local id.

-   -   If the primary id matches the global id (0x00, common to all QA         Chips), and the following byte from the master is the Trim Mode         id byte, and the fuse has not yet been blown, the QA Chip enters         Trim Mode and starts counting the number of internal clock         cycles until the next byte is received. Trim Mode cannot be         entered if the fuse has been blown.     -   If the primary id matches the global id (0x00, common to all QA         Chips), and the following byte from the master is the Program         Mode id byte, and the fuse has not yet been blown, the QA Chip         enters Program Mode. Program Mode cannot be entered if the fuse         has been blown.     -   If the primary id matches the global id (0x00, common to all QA         Chips), and the following byte from the master is the Active         Mode id bytes, the QA Chip enters Active Mode and executes         startup code, allowing the chip to set itself into a state to         subsequently receive authentication commands (includes setting a         local id and a trim value).     -   If the primary id matches the chip's local ID, the QA Chip         enters Active Mode, allowing the subsequent command to be         executed.

The valid 8-bit serial mode values sent after a global id are as shown in Table 349:

TABLE 349 Command byte values to place chip in specific mode Value Interpretation 10101011 Trim Mode (0xAB) (only functions when the fuse has not been blown) 10001101 Program Mode (0xAD) (only functions when the fuse has not been blown) 00000110 Active Mode (0x06) (resets the chip & loads the localId) 5.1.3 Trim Mode

Trim Mode is enabled by sending a global id byte (0x00) followed by the Trim Mode command byte (1xAB). Trim Mode can only be entered while the fuse has not yet been blown.

The purpose of Trim Mode is to set the trim value (an internal register setting) of the internal ring oscillator so that Flash erasures and writes are of the correct duration. This is necessary due to the 2:1 variation of the clock speed due to process variations. If writes an erasures are too long, the Flash memory will wear out faster than desired, and in some cases can even be damaged. Note that the 2:1 variation due to temperature still remains, so the effective operating speed of the chip is 7-14 MHz around a nominal 10 MHz.

Trim Mode works by measuring the number of system clock cycles that occur inside the chip from the receipt of the Trim Mode command byte until the receipt of a data byte. When the data byte is received, the data byte is copied to the trim register and the current value of the count is transmitted to the outside world.

Once the count has been transmitted, the QA Chip returns to Idle Mode.

At reset, the internal trim register setting is set to a known value r. The external user can now perform the following operations:

-   -   send the global id+write followed by the Trim Mode command byte     -   send the 8-bit value v over a specified time t     -   send a stop bit to signify no more data     -   send the global id+read followed by the Trim Mode command byte     -   receive the count c     -   send a stop bit to signify no more data

At the end of this procedure, the trim register will be v, and the external user will know the relationship between external time t and internal time c. Therefore a new value for v can be calculated.

The Trim Mode procedure can be repeated a number of times, varying both t and v in known ways, measuring the resultant c. At the end of the process, the final value for v is established (and stored in the trim register for subsequent use in Program Mode). This value v must also be written to the flash for later use (every time the chip is placed in Active Mode for the first time after power-up). For more information about the internal workings of Trim Mode and the accuracy of trim in the QA Chip, see Section 11.2 on page 967.

5.1.4 Program Mode

Program Mode is enabled by sending a global id byte (0x00) followed by the Program Mode command byte.

If the QA Chip knows already that the fuse has been blown, it simply does not enter Program Mode. If the QA Chip does not know the state of the fuse, it determines whether or not the internal fuse has been blown by reading 32-bit word 0 of the information block of flash memory. If the fuse has been blown the remainder of data from the Program Mode command is ignored, and the QA Chip returns to Idle Mode.

If the fuse is still intact, the chip enters Program Mode and erases the entire contents of Flash memory. The QA Chip then validates the erasure. If the erasure was successful, the QA Chip receives up to 4096 bytes of data corresponding to the new program code and variable data. The bytes are transferred in order byte₀ to byte₄₀₉₅.

Once all bytes of data have been loaded into Flash, the QA Chip returns to Idle Mode.

Note that Trim Mode functionality must be performed before a chip enters Program Mode for the first time. Otherwise the erasure and write durations could be incorrect.

Once the desired number of bytes have been downloaded in Program Mode, the LSS Master must wait for 80 μs (the time taken to write two bytes to flash at nybble rates) before sending the new transaction (e.g. Active Mode). Otherwise the last nybbles may not be written to flash.

5.1.5 After Manufacture

Directly after manufacture the flash memory will be invalid and the fuse will not have been blown. Therefore power-on-reset will not cause Active Mode. Trim Mode must therefore be entered first, and only after a suitable trim value is found, should Program Mode be entered to store a program. Active Mode can be entered if the program is known to be valid.

Logical View of CPU

6 Introduction

The QA Chip is a 32-bit microprocessor with on-board RAM for scratch storage, on-board flash for program storage, a serial interface, and specific security enhancements.

The high level commands that a user of an QA Chip sees are all implemented as small programs written in the CPU instruction set.

The following sections describe the memory model, the various registers, and the instruction set of the CPU.

7 Memory Model

The QA Chip has its own internal memory, broken into the following conceptual regions:

-   -   RAM variables (3 Kbits=96 entries at 32-bits wide), used for         scratch storage (e.g. HMAC-SHA1processing).     -   Flash memory (8 Kbytes main block(+)128 bytes info block) used         to hold the non-volatile authentication variables (including         program keys etc), and program code. Only 4 KBytes(+)64 bytes is         visible to the program addressing space due to shadowing.         Shadowing is where half of each byte is used to validate and         verify the other half, thus protecting against certain forms of         physical and logical attacks. As a result, two bytes are read to         obtain a single byte of data (this happens transparently).         7.1 RAM

The RAM region consists of 96×32-bit words required for the general functioning of the QA Chip, but only during the operation of the chip. RAM is volatile memory: once power is removed, the values are lost. Note that in actual fact memory retains its value for some period of time after power-down, but cannot be considered to be available upon power-up. This has issues for security that are addressed in other sections of this document.

RAM is typically used for temporary storage of variables during chip operation. Short programs can also be stored and executed from the RAM.

RAM is addressed from 0 to 5F. Since RAM is in an unknown state upon a RESET (RstL), program code should not assume the contents to be 0. Program code can, however, set the RAM to be a particular known state during execution of the reset command (guaranteed to be received before any other commands).

7.2 Flash Variables

The flash memory region contains the non-volatile information in the QA Chip. Flash memory retains its value after a RESET or if power is removed, and can be expected to be unchanged when the power is next turned on.

Byte 0 of main memory is the first byte of the program run for the command dispatcher. Note that the command dispatcher is always run with shadows enabled.

Bytes 0-7 of the information block flash memory is reserved as follows:

-   -   byte 0-3=fuse. A value of 0x5555AAAA indicates that the fuse has         been blown (think of a physical fuse whose wire is no longer         intact).     -   bytes 4-7=random number used to XOR all data for RAM and flash         memory accesses

After power-on reset (when the fuse is blown) or upon receipt of a globalId Active command, the 32-bit data from bytes 4-7 in the information block of Flash memory is loaded into an internal ChipMask register. In Active Mode (the chip is executing program code), all data read from the flash and RAM is XORed with the ChipMask register, and all data written to the flash and RAM is XORed with the ChipMask register before being written out. This XORing happens completely transparently to the program code. Main flash memory byte 0 onward is the start of program code. Note that byte 0 onward needs to be valid after being XORed with the appropriate bytes of ChipMask.

Even though CPU access is in 8-bit and 32-bit quantities, the data is actually stored in flash a nybble-at-a-time. Each nybble write is written as a byte containing 4 sets of b/

b pairs. Thus every byte write to flash is writing a nybble to real and shadow. A write mask allows the individual targetting of nybble-at-a-time writes.

The checking of flash vs shadow flash is automatically carried out each read (each byte contains both flash and shadow flash). If all 8 bits are 1, the byte is considered to be in its erased form¹, and returns 0 as the nybble. Otherwise, the value returned for the nybble depends on the size of the overall access and the setting of bit 0 of the 8-bit WriteMask. ¹TSMC's flash memory has an erased state of all 1s

-   -   All 8-bit accesses (i.e. instruction and program code fetches)         are checked to ensure that each byte read from flash is 4 sets         of b/         b pairs. If the data is not of this form, the chip hangs until a         new command is issued over the serial interface.     -   With 32-bit accesses (i.e. data used by program code), each byte         read from flash is checked to ensure that it is 4 sets of b/         b pairs. A setting of WriteMask₀=0 means that if the data is not         valID, then the chip will hang until a new command is issued         over the serial interface. A setting of WriteMask₀=1 means that         each invalid nybble is replaced by the upper nybble of the         WriteMask. This allows recovery after a write or erasure is         interrupted by a power-down.         8 Registers

A number of registers are defined for use by the CPU. They are used for control, temporary storage, arithmetic functions, counting and indexing, and for I/O.

These registers do not need to be kept in non-volatile (Flash) memory. They can be read or written without the need for an erase cycle (unlike Flash memory). Temporary storage registers that contain secret information still need to be protected from physical attack by Tamper Prevention and Detection circuitry and parity checks.

All registers are cleared to 0 on a RESET. However, program code should not assume any RAM contents have any particular state, and should set up register values appropriately. In particular, at the startup entry point, the various address registers need to be set up from unknown states.

8.1 GO

A 1-bit GO register is 1 when the program is executing, and 0 when it is not. Programs can clear the GO register to halt execution of program code once the command has finished executing.

8.2 Accumulator and Z Flag

The Accumulator is a 32-bit general-purpose register that can be thought of as the single data register. It is used as one of the inputs to all arithmetic operations, and is the register used for transferring information between memory registers.

The Z register is a 1-bit flag, and is updated each time the Accumulator is written to. The Z register contains the zero-ness of the Accumulator. Z=1 if the last value written to the Accumulator was 0, and 0 if the last value written was non-0.

Both the Accumulator and Z registers are directly accessible from the instruction set.

8.3 Address Registers

8.3.1 Program Counter Array and Stack Pointer

A 12-level deep 12-bit Program Counter Array (PCA) is defined. It is indexed by a 4-bit Stack Pointer (SP). The current Program Counter (PC), containing the address of the currently executing instruction, is effectively PCA[SP]. A single register bit, PCRamSel determines whether the program is executing from flash or RAM (0=flash, 1=RAM).

The PC is affected by calling subroutines or returning from them, and by executing branching instructions. The SP is affected by calling subroutines or returning from them. There is no bounds checking on calling too many subroutines: the oldest entry in the execution stack will be lost.

The entry point for program code is defined to be address 0 in Flash. This entry point is used whenever the master signals a new transaction.

8.3.2 A0-A3

There are 4 8-bit address registers Each register has an associated memory mode bit designating the address as in Flash (0) or RAM (1).

When an An register is pointing to an address in RAM, it holds the word number. When it is pointing to an address in Flash, it points to a set of 32-bit words that start at a 128-bit (16 byte) alignment. The A0 register has a special use of direct offset e.g. access is possible to (A0),0-7 which is the 32-bit word pointed to by A0 offset by the specified number of words.

8.3.3 WriteMask

The WriteMask register is used to determine how many nybbles will be written during a 32-bit write to Flash, and whether or not an invalid nybble will be replaced during a read from Flash.

During writes to flash, bit n (of 8) determines whether nybble n is written. The unit of writing is a nybble since half of each byte is used for shadow data. A setting of 0xFF means that all 32-bits will be written to flash (as 8 sets of nybble writes).

During 32-bit reads from flash (occurs as 8 reads), the value of WriteMask₀ is used to determine whether a read of invalid data is replaced by the upper nybble of WriteMask. If 0, a read of invalid data is not replaced, and the chip hangs until a new command is issued over the serial interface. If 1, a read of invalid data is replaced by the upper nybble of the WriteMask.

Thus a WriteMask setting of 0 (reset setting) means that no writes will occur to flash, and all reads are not replaced (causing the program to hang if an invalid value is encountered).

8.4 Counters

A number of special purpose counters/index registers are defined:

TABLE 350 Counter/Index registers Register Name Size Bits Description C1 1 × 3 3 Counter used to index arrays and general purpose counter C2 1 × 6 6 General purpose counter and can be used to index arrays

All these counter registers are directly accessible from the instruction set. Special instructions exist to load them with specific values, and other instructions exist to decrement or increment them, or to branch depending on the whether or not the specific counter is zero.

There are also 2 special flags (not registers) associated with C1 and C2, and these flags hold the zero-ness of C1 or C2. The flags are used for loop control, and are listed here, for although they are not registers, they can be tested like registers.

TABLE 351 Flags for testing C1 and C2 Name Description C1Z 1 = C1 is current zero, 0 = C1 is currently non-zero. C2Z 1 = C2 is current zero, 0 = C2 is currently non-zero. 8.5 RTMP

The single bit register RTMP allows the implementation of LFSRs and multiple precision shift registers.

During a rotate right (ROR) instruction with operand of RB, the bit shifted out (formally bit 0) is written to the RTMP register. The bit currently in the RTMP register becomes the new bit 31 of the Accumulator. Performing multiple ROR RB commands over several 32-bit values implements a multiple precision rotate/shift right.

The XRB operand operates in the same way as RB, in that the current value in the RTMP register becomes the new bit 31 of the Accumulator. However with the XRB instruction, the bit formally known as bit 0 does not simply replace RTMP (as in the RB instruction). Instead, it is XORed with RTMP, and the result stored in RTMP, thereby allowing the implementation of long LFSRs.

8.6 Registers Used for I/O

Several registers are defined for communication between the master and the QA Chip. These registers are LocalID, InByte and OutByte.

LocalId (7 bits) defines the chip-specific id that this particular QA Chip will accept commands for.

InByte (8 bits) provides the means for the QA Chip to obtain the next byte from the master. OutByte (8 bits) provides the means for the QA Chip to send a byte of data to the master.

From the QA Chip's point of view:

-   -   Reads from InByte will hang until there is 1 byte of data         present from the master.     -   Writes to OutByte will hang if the master has not already         consumed the last OutByte.

When the master begins a new command transaction, any existing data in InByte and OutByte is lost, and the PC is reset to the entry point in the code, thus ensuring correct framing of data.

8.7 Registers Used for Trimming Clock Speed

A single 8-bit Trim register is used to trim the ring oscillaor clock speed. The register has a known value of 0x00 during reset to ensure that reads from flash will succeed at the fastest process corners, and can be set in one of two ways:

-   -   via Trim Mode, which is necessary before the QA Chip is         programmed for the first time; or     -   via the CPU, which is necessary every time the QA Chip is         powered up before any flash write or erasure accesses can be         carried out.         8.8 Registers Used for Testing Flash

There are a number of registers specifically for testing the flash implementation. A single 32-bit write to an appropriate RAM address allows the setting of any combination of these flash test registers.

RAM consists of 96×32-bit words, and can be pointed to by any of the standard An address registers. A write to a RAM address in the range 97-127 does nothing with the RAM (reads return 0), but a write to a RAM address in the range 0x80-0x87 will write to specific groupings of registers according to the low 3 bits of the RAM address. A 1 in the address bit means the appropriate part of the 32-bit Accumulator value will be written to the appropriate flash test registers. A 0 in the address bit means the register bits will be unaffected.

The registers and address bit groupings are listed in Table 352:

TABLE 352 Flash test registers settable from CPU in RAM address range 0x80-0x87² adr data bitSuperscriptparanumonly bits name description 0 0 shadowsOff 0 = shadowing applies (nybble based flash access) 1 = shadowing disabled, 8-bit direct accesses to flash. 1 hiFlashAdr Only valid when shadowsOff = 1 0 = accesses are to lower 4 Kbytes of flash 1 = accesses are to upper 4 Kbytes of flash 2 1 3 enableFlashTest 0 = keep flash test register within the TSMC flash IP in its reset state 1 = enable flash test register to take on non-reset values. 8-4 flashTest Internal 5-bit flash test register within the TSMC flash IP (SFC008_08B9_HE). If this is written with 0x1E, then subsequent writes will be according to the TSMC write test mode. You must write a non-0x1E value or reset the register to exit this mode. 2 28-9  flashTime When timerSel is 1, this value is used for the duration of the program cycle within a standard flash write or erasure. 1 unit = 16 clock cycles (16 × 100 ns typical). Regardless of timerSel, this value is also used for the timeout following power down detection before the QA Chip resets itself. 1 unit = 1 clock cycle (=100 ns typical). Note that this means the programmer should set this to an appropriate value (e.g. 5 μs), just as the localId needs to be set. 29 timerSel 0 = use internal (default) timings for flash writes & erasures 1 = use flashTime for flash writes and erasures ²This is from the programmer's perspective. Addresses sent from the CPU are byte aligned, so the MRU needs to test bit n + 2. Similarly, checking DRAM address > 128 means testing bit 7 of the address in the CPU, and bit 9 in the MRU. ³unshadowed ⁴shadowed

When none of the address register bits 0-2 are set (e.g. a write to RAM address 0x80), then invalid writes will clear the illChip and retryCount registers.

For example, set the A0 register to be 0x80 in RAM. A write to (A0),0 will write to none of the flash test registers, but will clear the illChip and retryCount registers. A write to (A0),7 will write to all of the flash test registers. A write to (A0),2 will write to the enableFlashTest and flashTest registers only. A write to (A0),4 will write to the flashTime and bmerSel registers etc.

Finally, a write to address 0x88 in RAM will cause a device erasure. If infoBlockSel is 0, then the device erasure will only be of main memory. If infoBlockSel is 1, then the device erasure is of both main memory and the information block (which will also clear the ChipMask and the Fuse).

Reads of invalid RAM areas will reveal information as follows:

-   -   all invalid addresses in RAM (e.g. 0x80) will return the illChip         flag in the low bit (illChip is set whenever 16 consecutive bad         reads occur for a single byte in memory)     -   all invalid addresses in RAM with the low address bit set (e.g.         0x81, or (A0),1 when A0 holds 0x80), will additionally return         the most recent retryCount setting (only updated by the chip         when a bad read occurs). i.e. bit 0=illChip, bits         4-1=retryCount.         8.9 Register Summary

Table 353 provides a summary of the registers used in the CPU.

TABLE 353 Register summary Register name Description #bits A[0-3] address registers 49 = 36 Acc Accumulator 32 C1 general purpose counter and index 3 C2 general purpose counter and index 6 IllChip gets set whenever more than 15 1 consecutive bad reads from flash occurred (and any program executing has hung) InByte input byte from outside world 8 Go determines whether CPU is 1 executing LocalId determines id for this 7 chip's IO OutByte output byte to outside world 8 Z zero flag for last xfer to Acc 1 PCA program counter array 1212 = 144  PCRamSel Program code is executing in 1 flash (0) or ram (1) RetryCount counts the number of retries 4 for bad reads RTMP bit used to alow multi-word 1 rotations SP stack pointer into PCA 4 Trim rims ring oscillator frequency 8 flash test various registers in the 30 registers embedded flash and flash access logic specifically for testing the flash memory TOTAL (bits) 295 8.10 Startup

Whenever the chip is powered up, or receives a ‘write’ command over the serial interface, the PC and PCRamSel get set to 0 and execution begins at 0 in Flash memory. The program (starting at 0) needs to determine how the program was started by reading the InByte register.

If the first byte read is 0xFF, the chip is being requested to perform software reset tasks. Execution of software reset can only be interrupted by a power down. The reset tasks include setting up RAM to contain known startup state information, setting up Trim and localID registers etc. The CPU signals that it is now ready to receive commands from an external device by writing to the OutByte register.

An external Master is able to read the OutByte (and any further outbytes that the CPU decides to send) if it so wishes by a read using the localId.

Otherwise the first byte read will be of the form where the least significant bit is 0, and bits 7-1 contain the localId of the device as read over the serial interface. This byte is usually discarded since it nominally only has a value of differentiation against a software reset request. The second and subsequent bytes contain the data message of a write using the localId. The CPU can prevent interruption during execution by writing 0 to the localId and then restoring the desired localId at the later stage.

9 Instruction Set

The CPU operates on 8-bit instructions and typically on 32-bit data items. Each instruction typically consists of an opcode and operand, although the number of bits allocated to opcode and operand varies between instructions.

9.1 Basic Opcodes (Summary)

The opcodes are summarized in Table 354:

Table 354. opcode bit pattern map

TABLE 354 Opcode bit pattern map Opcode Mnemonic Simple Description 0000xxxx JMP Jump 0001xxxx JSR Jump subroutine 0010xxxx TBR Test and branch 0011xxxx DBR Decrement and branch 0100xxxx SC Set counter to a value 0101xxxx ST Store Accumulator in specified location 0110000x — reserved 01100010 JPZ Jump to 0 01100011 JPI Jump indirect 011001xx — reserved 01101xxx — reserved 01110000 — reserved 01110001 ERA Erase page of flash memory pointed to by Accumulator 01110010 JSZ Jump to subroutine at at 0 01110011 JSI Jump subroutine indirect 01110100 RTS Return from subroutine 01110101 HALT Stop the CPU 0111011x — reserved 01111xxx LIA Load immediate value into address register 10000xxx AND Bitwise AND Accumulator 10001xxx OR Bitwise OR Accumulator 1001xxxx XOR Exclusive-OR Accumulator 1010xxxx ADD Add a 32 bit value to the Accumulator 1011xxxx LD Load Accumulator 1100xxxx ROR Rotate Accumulator right 11010xxx AND Bitwise AND Accumulator⁵ 11011xxx OR Bitwise OR Accumulator^(Superscriptparanumonly) 11100xxx XOR Bitwise XOR Accumulator^(Superscriptparanumonly) 11101xxx ADD Add a 32 bit value to the Accumulator^(Superscriptparanumonly) 11110xxx LD Load Accumulator^(Superscriptparanumonly) 11111xxx RIA Rotate Accumulator into address register ⁵immediate form of instruction

Table 355 is a summary of valid operands for each opcode. The table is ordered alphabetically by opcode mnemonic. The binary value for each operand can be found in the subsequent sections.

TABLE 355 Valid operands for opcodes Opcode Valid operands ADD immediate value (A0), offset (An), {C1, C2} [where n = 0-3] AND immediate value (A0), offset DBR {C1, C2}, offset ERA HALT JMP address JPI JPZ JSI JSR address JSZ LIA {Flash, Ram}, An [where n = 0-3], {immediate value} LD immediate value (A0), offset (An), {C1, C2} [where n = 0-3] OR immediate value (A0), offset RIA {Flash, Ram}, An [where n = 0-3] ROR {InByte, OutByte, WriteMask, ID, C1, C2, RB, XRB, 1, 3, 8, 24, 31} RTS SC {C1, C2}, {immediate value} ST (A0), offset (An), {C1, C2} [where n = 0-3] TBR {0, 1}, offset XOR immediate value (A0), offset (An), {C1, C2} [where n = 0-3]

Additional pseduo-opcodes (for programming convenience) are as follows:

-   -   DEC=ADD 0xFF..     -   INC=ADD 0x01     -   NOT=XOR 0xFF..     -   LDZ=LD0     -   SC {C1, C2}, Acc=ROR {C1, C2}     -   RD=ROR Inbyte     -   WR=ROR OutByte     -   LDMASK=ROR WriteMask     -   LDID=ROR Id     -   NOP=XOR 0         9.2 Addressing Modes

The CPU supports a set of addressing modes as follows:

-   -   immediate     -   accumulator indirect     -   indirect fixed     -   indirect indexed         9.2.1 Immediate

In this form of addressing, the operand itself supplies the 32-bit data.

Immediate addressing relies on 3 bits of operand, plus an optional 8 bits at PC+1 to determine an 8-bit base value. Bits 0 to 1 of the opcode byte determine whether the base value comes from the opcode byte itself, or from PC+1, as shown in Table 356.

TABLE 356 Selection for base value in immediate mode Opcode¹⁻⁰ Base value 00 00000000 01 00000001 10 From PC + 1 (i.e. MIUData⁷⁻⁰) 11 11111111

The base value is computed by using CMD₀ as bit 0, and copying CMD₁ into the upper 7 bits.

The resultant 8 bit base value is then used as a 32-bit value, with 0s in the upper 24 bits, or the 8-bit value is replicated into the upper 32 bits. The selection is determined by bit 2 of the opcode byte, as follows:

TABLE 357 Replicate bits selection Opcode₂ Data 0 No replication. Data has 0 in upper 24 bits and baseVal in lower 8 bits 1 Replicated. Data is 32-bit value formed by replicating baseVal.

Opcodes that support immediate addressing are LD, ADD, XOR, AND, OR. The SC and LIA instructions are also immediate in that they store the data with the opcode, but they are not in the same form as that described here. See the detail on the individual instructions for more information. Single byte examples include:

-   -   LD 0     -   ADD 1     -   ADD 0xFF...# this subtracts 1 from the acc     -   XOR 0xFF...# this performs an effective logical NOT operation

Double byte examples include:

-   -   LD 0x05 # a constant     -   AND 0x0F # isolates the lower nybble     -   LD 0x36...# useful for HMAC processing         9.2.2 Accumulator Indirect

In this form of addressing, the Accumulator holds the effective address.

Opcodes that support Accumulator indirect addressing are JPI, JSI and ERA. In the case of JPI and JSI, the Accumulator holds the address to jump to. In the case of ERA, the Accumulator holds the address of the page in flash memory to be erased.

Examples include:

-   -   JPI     -   JSI     -   ERA         9.2.3 Indirect Fixed

In this form of addressing, address register A0 is used as a base address, and then a specific fixed offset is added to the base address to give the effective address.

Bits 2-0 of the opcode byte specify the fixed offset from A0, which means the fixed offset has a range of 0 to 7.

Opcodes that support indirect indexed addressing are LD, ST, ADD, XOR, AND, OR.

Examples include:

-   -   LD (A0),2     -   ADD (A0), 3     -   AND (A0), 4     -   ST (A0), 7         9.2.4 Indirect Indexed

In this form of addressing, an address register is used as a base address, and then an index register is used to offset from that base address to give the effective address.

The address register is one of 4, and is selected via bits 2-1 of the opcode byte as follows:

TABLE 358 Address register selection address register Opcode²⁻¹ selected 00 A0 01 A1 10 A2 11 A3

Bit 0 of the opcode byte selects whether index register C1 or C2 is used:

The counter is selected as follows:

TABLE 359 Interpretation of counter for DBR Opcode₀ interpretion 0 C1 1 C2

Opcodes that support indirect indexed addressing are LD, ST, ADD, XOR.

Examples include:

-   -   LD (A2), C1     -   ADD (A1), C1     -   ST (A3), C2

Since C1 and C2 can only decement, processing of data structures typically works by loading Cn with some number n and decrementing to 0. Thus (Ax),n is the first word accessed, and (Ax),0 is the last 32-bit word accessed in the loop.

9.3 ADD—Add to Accumulator

-   -   Mnemonic: ADD     -   Opcode: 1010xxxx, and 11101xxx     -   Usage: ADD effective-address, or ADD immediate-value

The ADD instruction adds the specified 32-bit value to the Accumulator via modulo 2³² addition.

The 11101xxx form of the opcode follows the immediate addressing rules (see Section 9.2.1 on page 946). The 1010xxxx form of the opcode defines an effective address as follows:

TABLE 360 Interpretation of operand for ADD (1010xxxx) bit 3 interpretion comment 0 (A0), offset indirect fixed addressing (see Section 9.2.3 on page 948) 1 (An), Cn indirect indexed addressing (see Section 9.2.4 on page 948)

The Z flag is also set during this operation, depending on whether the result (loaded into the Accumulator) is zero or not.

9.4 AND—Bitwise AND

-   -   Mnemonic: AND     -   Opcode: 10000xxx, and 11010xxx     -   Usage: AND effective-address, or AND immediate-value

The AND instruction performs a 32-bit bitwise AND operation on the Accumulator.

The 11010xxx form of the opcode follows the immediate addressing rules (see Section 9.2.1 on page 946). The 10000xxx form of the opcode follows the indirect fixed addressing rules (see Section 9.2.3 on page 948).

The Z flag is also set during this operation, depending on whether the resultant 32-bit value (loaded into the Accumulator) is zero or not.

9.5 DBR—Decrement and Branch

-   -   Mnemonic: DBR     -   OpCode: 0011xxxx     -   Usage: DBR Counter, Offset

This instruction provides the mechanism for building simple loops.

The counter is selected from bit 0 of the opcode byte as follows:

TABLE 361 Interpretation of counter for DBR bit 0 interpretion 0 C1 1 C2

If the specified counter is non-zero, then the counter is decremented and the designated offset is added to the current instruction address (PC for 1-byte instructions, PC+1 for 2-byte instructions). If the specified counter is zero, it is decremented (all bits in the counter become set) and processing continues at the next instruction (PC+1 or PC+2). The designated offset will typically be negative for use in loops.

The instruction is either 1 or two bytes, as determined by bits 3-1 of the opcode byte:

-   -   If bits 3-1=000, the instruction consumes 2 bytes. The 8 bits at         PC+1 are treated as a signed number and used as the offset         amount. Thus 0xFF is treated as −1, and 0x01 is treated as +1.     -   If bits 3-1≠000, the instruction consumes 1 byte. Bits 3-1 are         treated as a negative number (the sign bit is implied) and used         as the offset amount. Thus 111 is treated as −1, and 001 is         treated as −7. This is useful for small loops.

The effect is that if the branch is back 1-7 bytes (1 byte is not particularly useful), then the single byte form of the instruction can be used. If the branch is forward, or backward more than 7 bytes, then the 2-byte instruction is required.

9.6 ERA—Erase

-   -   Mnemonic: ERA     -   Opcode: 01110001     -   Usage: ERA

This instruction causes an erasure of the 256-byte page of flash memory pointed to by the Accumulator. The Accumulator is assumed to contain an 8-bit pointer to a 128-bit (16 byte) aligned structure (same structure as the address registers). The page number to be erased comes from bits 7-4, and the lower 4 bits are ignored.

Note that the size of the flash memory page being erased is actually 512 bytes, but in terms of data storage and addressing from the point of view of the CPU, there is only 256 bytes in the page.

9.7 HALT—Halt CPU Operation

-   -   Mnemonic: HALT     -   Opcode: 01110101     -   Usage: HALT

The HALT instruction writes a 0 to the internal GO register, thereby causing the CPU to terminate the currently executing program. The CPU will only be restarted with a new localId transaction from the Master or by a globalId plus Active Mode byte.

9.8 JMP—Jump

-   -   Mnemonic: JMP     -   Opcode: 0000xxxx     -   Usage: JMP effective-address

The JMP instruction provides for a method of branching to a specified address. The instruction loads the PC with the effective address.

The new PC is loaded as follows: bits 11-8 are obtained from bits 3-0 of the JMP opcode byte, and bits 7-0 are obtained from PC+1.

9.9 JPI—Jump Indirect

-   -   Mnemonic: JPI     -   Opcode: 01100011     -   Usage: JPI

The JPI instruction loads the PC with the lower 12 bits of the Accumulator, and sets the PCRamSel register with bit 15 of the Accumulator. Note that the stack is unaffected (unlike JSI).

9.10 JPZ—Jump to Zero

-   -   Mnemonic: JPZ     -   Opcode: 01100010     -   Usage: JPZ

The JPZ instruction loads the PC and PCRamSel with 0, thereby causing a jump to address 0 in Flash memory.

Programmers will not typically use the JPZ command. However the CPU executes this instruction whenever a new command arrives over the serial interface, so that the code entry point is known i.e. every time the chip receives a new command, execution begins at address 0 in flash. This does not change the status of any other internal register settings (e.g. the flash test registers).

9.11 JSI—Jump Subroutine Indirect

-   -   Mnemonic: JSI     -   Opcode: 01110011     -   Usage: JsI

The JSI instruction allows the jumping to a subroutine whose address is obtained from the Accumulator. The instruction pushes the current PC onto the stack, loads the PC with the lower 12 bits of the Accumulator, and sets the PCRamSel register with bit 15 of the Accumulator.

The stack provides for 12 levels of execution (11 subroutines deep). It is the responsibility of the programmer to ensure that this depth is not exceeded or the deepest return value will be overwritten (since the stack wraps). Programs can take advantage of the fact that the stack wraps.

9.12 JSR—Jump Subroutine

-   -   Mnemonic: JSR     -   Opcode: 0001xxxx     -   Usage: JSR effective-address

The JSR instruction provides for the most common usage of the subroutine construct. The instruction pushes the current PC onto the stack, and loads the PC with the effective address.

The new PC is loaded as follows: bits 11-8 are obtained from bits 3-0 of the JSR opcode byte, and bits 7-0 are obtained from PC+1.

The stack provides for 12 levels of execution (11 subroutines deep). It is the responsibility of the programmer to ensure that this depth is not exceeded or the return value will be overwritten (since the stack wraps). Programs can take advantage of the fact that the stack wraps.

9.13 JSZ—Jump to Subroutine at Zero

-   -   Mnemonic: JSZ     -   Opcode: 01110010     -   Usage: JSZ

The JSZ instruction jumps to the subroutine at flash address 0 (i.e. it pushes the current PC onto the stack, and loads the PC and PCRamSel with 0).

Programmers will not typically use the JSZ command. It exists merely as a result of opcode decoding minimization and can be used to assist with the testing of the chip.

9.14 LD—Load Accumulator

-   -   Mnemonic: LD     -   Opcode: 1011xxxx, and 11110xxx         -   Usage: LD effective-address, or LD immediate-value

The LD instruction loads the Accumulator with the 32-bit value.

The 11110xxx form of the opcode follows the immediate addressing rules (see Section 9.2.1 on page 946). The 1011xxxx form of the opcode defines an effective address as follows:

TABLE 362 Interpretation of operand for LD (1011xxxx) bit 3 interpretion comment 0 (A0), offset indirect fixed addressing (see Section 9.2.3 on page 948) 1 (An), Cn indirect indexed addressing (see Section 9.2.4 on page 948)

The Z flag is also set during this operation, depending on whether the value loaded into the Accumulator is zero or not.

9.15 LIA—Load Immediate Address

-   -   Mnemonic: LIA     -   Opoode: 01111xxx     -   Usage: LIAF AddressRegister, Value # for flash addresses         -   LIAR AddressRegister, Value # for ram addresses

The LIA instruction transfers the data from PC+1 into the designated address register (A0-A3), and sets the memory mode bit for that address register.

Bit 0 specifies whether the address is in flash or ram, as follows:

TABLE 363 Interpretation of memory mode for LIA bit 0 interpretion 0 Flash 1 Ram

The address register to be targetted is selected via bits 2-1 of the instruction.

9.16 OR—Bitwise OR

-   -   Mnemonic: OR     -   Opcode: 10001xxx, and 11011xxx     -   Usage: OR effective-address, or OR immediate-value

The OR instruction performs a 32-bit bitwise OR operation on the Accumulator.

The 11011xxx form of the opcode follows the immediate addressing rules (see Section 9.2.1 on page 946). The 10001xxx form of the opcode follows the indirect fixed addressing rules (see Section 9.2.3 on page 948).

The Z flag is also set during this operation, depending on whether the resultant 32-bit value (loaded into the Accumulator) is zero or not.

9.17 RIA—Rotate in Address

-   -   Mnemonic: RIA     -   Opcode: 11111xxx     -   Usage: RIAF AddressRegister # for flash addresses         -   RIAR AddressRegister # for ram addresses

The RIA instruction transfers the lower 8 bits of the Accumulator into the designated address register (A0-A3), sets the memory mode bit for that address register, and rotates the Accumulator right by 8 bits.

Bit 0 specifies whether the address is in flash or ram, as follows:

TABLE 364 Interpretation of memory mode for RIA bit 0 interpretion 0 Flash 1 Ram

The address register to be targetted is selected via bits 2-1 of the instruction.

9.18 ROR—Rotate Right

-   -   Mnemonic: ROR     -   Opcode: 1100xxxx     -   Usage: ROR Value

The ROR instruction provides a way of rotating the Accumulator right a set number of bits. The bit(s) coming in at the top of the Accumulator (to become bit 31) can either come from the previous lower bits of the Accumulator, from the serial connection, or from external flags. The bit(s) rotated out can also be output from the serial connection, or combined with an external flag.

The allowed operands are as follows:

TABLE 365 Interpretation of operand for ROR bits 3-0 interpretion 0000 RB 0001 XRB 0010 WriteMask 0011 1 0100 - (reserved) 0101 3 0110 31 0111 24 1000 C1 1001 C2 1010 - (reserved) 1011 - (reserved) 1100 8 1101 ID 1110 InByte 1111 OutByte

The Z flag is also set during this operation, depending on whether resultant 32-bit value (loaded into the Accumulator) is zero or not.

In its simplest form, the operand for the ROR instruction is one of 1, 3, 8, 24, 31, indicating how many bit positions the Accumulator should be rotated. For these operands, there is no external input or output—the bits of the Accumulator are merely rotated right. Note that these values are the equivalent to rotating left 31, 29, 24, 8, 1 bit positions.

With operand WriteMask, the lower 8 bits of the Accumulator are transferred to the WriteMask register, and the Accumulator is rotated right by 1 bit. This conveniently allows successive nybbles to be masked during Flash writes if the Accumulator has been preloaded with an appropriate value (eg 1x01).

With operands C1 and C2, the lower appropriate number of bits of the Accumulator (3 for C1, 6 for C2) are transferred to the C1 or C2 register and the lower 6 bits of the Accumulator are loaded with the previous value of the Cn register. The remaining upper bits of the Accumulator are set as follows: bit 31-24 are copied from previous bits 7-0, and bits 23-6 are copied from previous bits 31-14 (effectively junk). As a result, the Accumulator should be subsequently masked if the programmer wants to compare for specific values).

With operand ID, the 7 low-order bits are transferred from the Accumulator to the LocalId register, the low-order 8 bits of the Accumulator are copied to the Trim register if the Trim register has not already been written to after power-on reset, and the Accumulator is rotated right by 8 bits. This means that the ROR ID instruction needs to be performed twice, typically during Global Active Mode—once to set Trim, and once to set LocalId. Note there is no way to read the contents of the localId or Trim registers directly. However the LocalId sent to the program for a command is available as bits 7-1 of the first byte obtained from InByte after program startup.

With operand InByte, the next serial input byte is transferred to the highest 8 bits of the Accumulator. The InByteValid bit is also cleared. If there is no input byte available from the client yet, execution is suspended until there is one. The remainder of the Accumulator is shifted right 8 bit positions (bit 31 becomes bit 23 etc.), with lowest bits of the Accumulator shifted out.

With operand OutByte, the Accumulator is shifted right 8 bit positions. The byte shifted out from bits 7-0 is stored in the OutByte register and the OutByteValid flag is set. It is therefore ready for a client to read. If the OutByteValid flag is already set, execution of the instruction stalls until the OutByteValid flag cleared (when the OutByte byte has been read by the client). The new data shifted in to the upper 8 bits of the Accumulator is what was transferred to the OutByte register (i.e. from the Accumulator).

Finally, the RB and XRB operands allow the implementation of LFSRs and multiple precision shift registers. With RB, the bit shifted out (formally bit 0) is written to the RTMP register. The register currently in the RTMP register becomes the new bit 31 of the Accumulator. Performing multiple ROR RB commands over several 32-bit values implements a multiple precision rotate/shift right. The XRB operates in the same way as RB, in that the current value in the RTMP register becomes the new bit 31 of the Accumulator. However with the XRB instruction, the bit formally known as bit 0 does not simply replace RTMP (as in the RB instruction). Instead, it is XORed with RTMP, and the result stored in RTMP. This allows the implementation of long LFSR_(S), as required by the authentication protocol.

9.19 RTS—Return From Subroutine

-   -   Mnemonic: RTS     -   Opcode: 01110100     -   Usage: RTS

The RTS instruction pulls the saved PC from the stack, adds 1, and resumes execution at the resultant address. The effect is to cause execution to resume at the instruction after the most recently executed JSR or JSI instruction.

Although 12 levels of execution are provided for (11 subroutines), it is the responsibility of the programmer to balance each JSR and JSI instruction with an RTS. A RTS executed with no previous JSR will cause execution to begin at whatever address happens to be pulled from the stack. Of course this may be desired behaviour in specific circumstances.

9.20 SC—Set Counter

-   -   Mnemonic: SC     -   Opcode: 0100xxxx     -   Usage: SC Counter Value

The SC instruction is used to transfer a 3-bit Value into the specified counter. The operand determines which of counters C1 and C2 is to be loaded as well as the value to be loaded. Value is stored in bits 3-1 of the 8-bit opcode, and the counter is specified by bit 0 as follows:

TABLE 366 Interpretation of counter for SC bit 0 interpretion 0 C1 1 C2

Since counter C1 is 3 bits, Value is copied directly into C1.

For counter C2, C2₂₋₀ are copied to C2₅₋₃, and Value is copied to C2₂₋₀. Two SC C2 instructions are therefore required to load C2 with a given 6-bit value. For example, to load C2 with 0x0C, we would have SC C2 1 followed by SC C2 4.

9.21 ST—Store Accumulator

-   -   Mnemonic: ST     -   Opcode: 0101xxxx     -   Usage: ST effective-address

The ST instruction stores the 32-bit Accumulator at the effective address. The effective address is determined as follows:

TABLE 367 Interpretation of operand for ST (0101xxxx) bit 3 interpretion comment 0 (A0), offset indirect fixed addressing (see Section 9.2.3 on page 948) 1 (An), Cn indirect indexed addressing (see Section 9.2.4 on page 948)

If the effective address in Flash memory, only those nybbles whose corresponding WriteMask bit is set will be written to Flash. Programmers should be very aware of flash characteristics (write time, longevity, page size etc. when storing data in flash).

There is always the possibility that power could be removed during a write to Flash. If this occurs, the flash will be in an indeterminate state. If the QA Chip is warned by the external system that power is about to be removed (via the master causing a transition to Idle Mode), the write will be aborted cleanly at the nearest nybble boundary (writes occur in the order of least significant to most significant).

9.22 TBR—Test and Branch

-   -   Mnemonic: TBR     -   Opcode: 0010xxxx     -   Usage: TBR Value Offset

The Test and Branch instruction tests the status of the Z flag (the zero-ness of the Accumulator), and then branches if a match ocurs.

The zero-ness is selected from bit 0 of the opcode byte as follows:

TABLE 368 Interpretation of zero-ness for TBR bit 0 interpretion 0 true if Acc is zero (Z = 1) 1 true if Acc is non-zero (Z = 0)

If the specified zero-test matches, then the designated offset is added to the current instruction address (PC for 1-byte instructions, PC+1 for 2-byte instructions). If the zero-test does not match, processing continues at the next instruction (PC+1 or PC+2). The instruction is either 1 or two bytes, as determined by bits 3-1 of the opcode byte:

-   -   If bits 3-1=000, the instruction consumes 2 bytes. The 8 bits at         PC+1 are treated as a signed number and used as the offset         amount to be added to PC+1. Thus 0xFF is treated as −1, and 0x01         is treated as +1.     -   If bits 3-1≠000, the instruction consumes 1 byte. Bits 3-1 are         treated as a positive number (the sign bit is implied) and used         as the offset amount to be added to PC. Thus 111 is treated as         7, and 001 is treated as 1. This is useful for skipping over a         small number of instructions.

The effect is that if the branch is forward 1-7 bytes (1 byte is not particularly useful), then the single byte form of the instruction can be used. If the branch is backward, or forward more than 7 bytes, then the 2-byte instruction is required.

9.23 XOR—Bitwise Exclusive OR

-   -   Mnemonic: XOR     -   Opcode: 1001xxxx, and 11100xxx     -   Usage: XOR effective-address, or XOR immediate-value

The XOR instruction performs a 32-bit bitwise XOR operation on the Accumulator.

The 11100xxx form of the opcode follows the immediate addressing rules (see Section 9.2.1 on page 946). The 1001xxxx form of the opcode has an effective address as follows:

TABLE 369 Interpretation of operand for XOR (1001xxxx) bit 3 interpretion comment 0 (A0), offset indirect fixed addressing (see Section 9.2.3 on page 948) 1 (An), Cn indirect indexed addressing (see Section 9.2.4 on page 948)

The Z flag is also set during this operation, depending on whether the result (loaded into the Accumulator) is zero or not.

Implementation

10 Introduction

This chapter provides the high-level definition of a CPU capable of implementing the functionality required of an QA Chip.

10.1 Physical Interface

10.1.1 Pin Connections

The pin connections are described in Table 370.

TABLE 370 Pin connections to QA Chip pin direction description Vdd In Nominal voltage. If the voltage deviates from this by more than a fixed amount, the chip will RESET. GND In SClk In Serial clock SDa In/Out Serial data

The system operating clock SysClk is different to SClk. SysClk is derived from an internal ring oscillator based on the process technology. In the FPGA implementation SysClk is obtained via a 5th pin.

10.1.2 Size and Cost

The QA Chip uses a 0.25 μm CMOS Flash process for an area of 1 mm² yielding a 10 cent manufacturing cost in 2002. A breakdown of area is listed in Table 371.

TABLE 371 Breakdown of Area for QA Chip approximate area (mm²) description 0.49 8 KByte flash memory TSMC: SFC0008_08B9_HE (8K × 8-bits, erase page size = 512 bytes) Area = 724.688 μm × 682.05 μm. 0.08 3072 bits of static RAM 0.38 General logic 0.05 Analog circuitry 1 TOTAL (approximate)

Note that there is no specific test circuitry (scan chains or BIST) within the QA Chip (see Section 10.3.10 on page 965), so the total transistor count is as shown in Table 371.

10.1.3 Reset

The chip performs a RESET upon power-up. In addition, tamper detection and prevention circuitry in the chip will cause the chip to either RESET or erase Flash memory (depending on the attack detected) if an attack is detected.

10.2 Operating Speed

The base operating system clock SysClk is generated internally from a ring oscillator (process dependant). Since the frequency varies with operating temperature and voltage, the clock is passed through a temperature-based clock filter before use (see Section 10.3.3 on page 961). The frequency is built into the chip during manufacture, and cannot be changed. The frequency is in the range 7-14 MHz.

10.3 General Manufacturing Comments

Manufacturing comments are not normally made when normally describing the architecture of a chip. However, in the case of the QA Chip, the physical implementation of the chip is very much tied to the security of the key. Consequently a number of specialized circuits and components are necessary for implementation of the QA Chip. They are listed here.

-   -   Flash process     -   Internal randomized clock     -   Temperature based clock filter     -   Noise generator     -   Tamper Prevention and Detection circuitry     -   Protected memory with tamper detection     -   Boot-strap circuitry for loading program code     -   Data connections in polysilicon layers where possible     -   OverUnderPower Detection Unit     -   No scan-chains or BIST         10.3.1 Flash Process

The QA Chip is implemented with a standard Flash manufacturing process. It is important that a Flash process be used to ensure that good endurance is achieved (parts of the Flash memory can be erased/written many times).

10.3.2 Internal Randomized Clock

To prevent clock glitching and external clock-based attacks, the operating clock of the chip should be generated internally. This can be conveniently accomplished by an internal ring oscillator. The length of the ring depends on the process used for manufacturing the chip.

Due to process and temperature variations, the clock needs to be trimmed to bring it into a range usable for timing of Flash memory writes and erases.

The internal clock should also contain a small amount of randomization to prevent attacks where light emissions from switching events are captured, as described below.

Finally, the generated clock must be passed through a temperature-based clock filter before being used by the rest of the chip (see Section 10.3.3 on page 961).

The normal situation for FET implementation for the case of a CMOS inverter (which involves a pMOS transistor combined with an nMOS transistor) as shown in FIG. 353.

During the transition, there is a small period of time where both the nMOS transistor and the pMOS transistor have an intermediate resistance. The resultant power-ground short circuit causes a temporary increase in the current, and in fact accounts for around 20% of current consumed by a CMOS device. A small amount of infrared light is emitted during the short circuit, and can be viewed through the silicon substrate (silicon is transparent to infrared light). A small amount of light is also emitted during the charging and discharging of the transistor gate capacitance and transmission line capacitance.

For circuitry that manipulates secret key information, such information must be kept hidden.

Fortunately, IBM's PICA system and LVP (laser voltage probe) both have a requirement for repeatability due to the fact that the photo emissions are extremely weak (one photon requires more than 10⁵ switching events). PICA requires around 10⁹ pases to build a picture of the optical waveform. Similarly the LVP requires multiple passes to ensure an adequate SNR.

Randomizing the clock stops repeatability (from the point of view of collecting information about the same position in time), and therefore reduces the possibility of this attack.

10.3.3 Temperature Based Clock Filter

The QA Chip circuitry is designed to operate within a specific clock speed range. Although the clock is generated by an internal ring oscillator, the speed varies with temperature and power. Since the user supplies the temperature and power, it is possible for an attacker to attempt to introduce race-conditions in the circuitry at specific times during processing. An example of this is where a low temperature causes a clock speed higher than the circuitry is designed for, and this may prevent an XOR from working properly, and of the two inputs, the first may always be returned. These styles of transient fault attacks are documented further in [1]. The lesson to be learned from this is that the input power and operating temperature cannot be trusted.

Since the chip contains a specific power filter, we must also filter the clock. This can be achieved with a temperature sensor that allows the clock pulses through only when the temperature range is such that the chip can function correctly.

The filtered clock signal would be further divided internally as required.

10.3.4 Noise Generator

Each QA Chip should contain a noise generator that generates continuous circuit noise. The noise will interfere with other electromagnetic emissions from the chip's regular activities and add noise to the I_(dd) signal. Placement of the noise generator is not an issue on an QA Chip due to the length of the emission wavelengths.

The noise generator is used to generate electronic noise, multiple state changes each clock cycle, and as a source of pseudo-random bits for the Tamper Prevention and Detection circuitry (see Section 10.3.5 on page 962).

A simple implementation of a noise generator is a 64-bit maximal period LFSR seeded with a non-zero number.

10.3.5 Tamper Prevention and Detection Circuitry

A set of circuits is required to test for and prevent physical attacks on the QA Chip. However what is actually detected as an attack may not be an intentional physical attack. It is therefore important to distinguish between these two types of attacks in an QA Chip:

-   -   where you can be certain that a physical attack has occurred.     -   where you cannot be certain that a physical attack has occurred.

The two types of detection differ in what is performed as a result of the detection. In the first case, where the circuitry can be certain that a true physical attack has occurred, erasure of flash memory key information is a sensible action. In the second case, where the circuitry cannot be sure if an attack has occurred, there is still certainly something wrong. Action must be taken, but the action should not be the erasure of secret key information. A suitable action to take in the second case is a chip RESET. If what was detected was an attack that has permanently damaged the chip, the same conditions will occur next time and the chip will RESET again. If, on the other hand, what was detected was part of the normal operating environment of the chip, a RESET will not harm the key.

A good example of an event that circuitry cannot have knowledge about, is a power glitch. The glitch may be an intentional attack, attempting to reveal information about the key. It may, however, be the result of a faulty connection, or simply the start of a power-down sequence. It is therefore best to only RESET the chip, and not erase the key. If the chip was powering down, nothing is lost.

If the System is faulty, repeated RESETs will cause the consumer to get the System repaired. In both cases the consumable is still intact.

A good example of an event that circuitry can have knowledge about, is the cutting of a data line within the chip. If this attack is somehow detected, it could only be a result of a faulty chip (manufacturing defect) or an attack. In either case, the erasure of the secret information is a sensible step to take.

Consequently each QA Chip should have 2 Tamper Detection Lines—one for definite attacks, and one for possible attacks. Connected to these Tamper Detection Lines would be a number of Tamper Detection test units, each testing for different forms of tampering. In addition, we want to ensure that the Tamper Detection Lines and Circuits themselves cannot also be tampered with.

At one end of the Tamper Detection Line is a source of pseudo-random bits (clocking at high speed compared to the general operating circuitry). The Noise Generator circuit described above is an adequate source. The generated bits pass through two different paths—one carries the original data, and the other carries the inverse of the data. The wires carrying these bits are in the layer above the general chip circuitry (for example, the memory, the key manipulation circuitry etc.). The wires must also cover the random bit generator. The bits are recombined at a number of places via an XOR gate. If the bits are different (they should be), a 1 is output, and used by the particular unit (for example, each output bit from a memory read should be ANDed with this bit value). The lines finally come together at the Flash memory Erase circuit, where a complete erasure is triggered by a 0 from the XOR. Attached to the line is a number of triggers, each detecting a physical attack on the chip. Each trigger has an oversize nMOS transistor attached to GND. The Tamper Detection Line physically goes through this nMOS transistor. If the test fails, the trigger causes the Tamper Detect Line to become 0. The XOR test will therefore fail on either this clock cycle or the next one (on average), thus RESETing or erasing the chip.

FIG. 349 illustrates the basic principle of a Tamper Detection Line in terms of tests and the XOR connected to either the Erase or RESET circuitry.

The Tamper Detection Line must go through the drain of an output transistor for each test, as illustrated by FIG. 350.

It is not possible to break the Tamper Detect Line since this would stop the flow of 1 s and 0s from the random source. The XOR tests would therefore fail. As the Tamper Detect Line physically passes through each test, it is not possible to eliminate any particular test without breaking the Tamper Detect Line.

It is important that the XORs take values from a variety of places along the Tamper Detect Lines in order to reduce the chances of an attack. FIG. 351 illustrates the taking of multiple XORs from the Tamper Detect Line to be used in the different parts of the chip. Each of these XORs can be considered to be generating a ChipOK bit that can be used within each unit or sub-unit.

A typical usage would be to have an OK bit in each unit that is ANDed with a given ChipOK bit each cycle. The OK bit is loaded with 1 on a RESET. If OK is 0, that unit will fail until the next RESET. If the Tamper Detect Line is functioning correctly, the chip will either RESET or erase all key information. If the RESET or erase circuitry has been destroyed, then this unit will not function, thus thwarting an attacker.

The destination of the RESET and Erase line and associated circuitry is very context sensitive. It needs to be protected in much the same way as the individual tamper tests. There is no point generating a RESET pulse if the attacker can simply cut the wire leading to the RESET circuitry.

The actual implementation will depend very much on what is to be cleared at RESET, and how those items are cleared.

Finally, FIG. 352 shows how the Tamper Lines cover the noise generator circuitry of the chip. The generator and NOT gate are on one level, while the Tamper Detect Lines run on a level above the generator.

10.3.6 Protected Memory with Tamper Detection

It is not enough to simply store secret information or program code in flash memory. The Flash memory and RAM must be protected from an attacker who would attempt to modify (or set) a particular bit of program code or key information. The mechanism used must conform to being used in the Tamper Detection Circuitry (described above).

The first part of the solution is to ensure that the Tamper Detection Line passes directly above each flash or RAM bit. This ensures that an attacker cannot probe the contents of flash or RAM. A breach of the covering wire is a break in the Tamper Detection Line. The breach causes the Erase signal to be set, thus deleting any contents of the memory. The high frequency noise on the Tamper Detection Line also obscures passive observation.

The second part of the solution for flash is to always store the data with its inverse. In each byte, 4 bits contains the data, and 4 bits (the shadow) contains the inverse of the data. If both are 0, this is a valid erase state, and the value is 0. Otherwise, the memory is only valid if the 4 bits of shadow are the inverse of the main 4 bits. The reasoning is that it is possible to add electrons to flash via a FIB, but not take electrons away. If it is possible to change a 0 to 1 for example, it is not possible to do the same to its inverse, and therefore regardless of the sense of flash, an attack can be detected.

The second part of the solution for RAM is to use a parity bit. The data part of the register can be checked against the parity bit (which will not match after an attack).

The bits coming from Flash and RAM can therefore be validated by a number of test units (one per bit) connected to the common Tamper Detection Line. The Tamper Detection circuitry would be the first circuitry the data passes through (thus stopping an attacker from cutting the data lines).

In addition, the data and program code should be stored in different locations for each chip, so an attacker does not know where to launch an attack. Finally, XORing the data coming in and going to Flash with a random number that varies for each chip means that the attacker cannot learn anything about the key by setting or clearing an individual bit that has a probability of being the key (the inverse of the key must also be stored somewhere in flash).

Finally, each time the chip is called, every flash location is read before performing any program code. This allows the flash tamper detection to be activated in a common spot instead of when the data is actually used or program code executed. This reduces the ability of an attacker to know exactly what was written to.

10.3.7 Boot-Strap Circuitry for Loading Program Code

Program code should be kept in protected flash instead of ROM, since ROM is subject to being altered in a non-testable way. A boot-strap mechanism is therefore required to load the program code into flash memory (flash memory is in an indeterminate state after manufacture).

The boot-strap circuitry must not be in a ROM—a small state-machine suffices. Otherwise the boot code could be trivially modified in an undetectable way.

The boot-strap circuitry must erase all flash memory, check to ensure the erasure worked, and then load the program code.

The program code should only be executed once the flash program memory has been validated via Program Mode.

Once the final program has been loaded, a fuse can be blown to prevent further programming of the chip.

10.3.8 Connections in Polysilicon Layers Where Possible

Wherever possible, the connections along which the key or secret data flows, should be made in the polysilicon layers. Where necessary, they can be in metal 1, but must never be in the top metal layer (containing the Tamper Detection Lines).

10.3.9 OverUnder Power Detection Unit

Each QA Chip requires an OverUnder Power Detection Unit (PDU) to prevent Power Supply Attacks. A PDU detects power glitches and tests the power level against a Voltage Reference to ensure it is within a certain tolerance. The Unit contains a single Voltage Reference and two comparators. The PDU would be connected into the RESET Tamper Detection Line, thus causing a RESET when triggered.

A side effect of the PDU is that as the voltage drops during a power-down, a RESET is triggered, thus erasing any work registers.

10.3.10 No Scan Chains or BIST

Test hardware on an QA Chip could very easily introduce vulnerabilities. In addition, due to the small size of the QA Chip logic, test hardware such as scan paths and BIST units could in fact take a sizeable chunk of the final chip, lowering yield and causing a situation where an error in the test hardware causes the chip to be unusable. As a result, the QA Chip should not contain any BIST or scan paths. Instead, the program memory must first be validated via the Program Mode mechanism, and then a series of program tests run to verify the remaining parts of the chip.

11 Architecture

FIG. 389 shows a high level block diagram of the QA Chip. Note that the tamper prevention and detection circuitry is not shown.

11.1 Analogue Unit

FIG. 390 shows a block diagram of the Analogue Unit. Blocks shown in yellow provide additional protection against physical and electrical attack and, depending on the level of security required, may optionally be implemented.

11.1.1 Ring Oscillator

The operating clock of the chip (SysClk) is generated by an internal ring oscillator whose frequency can be trimmed to reduce the variation from 4:1 (due to process and temperature) down to 2:1 (temperature variations only) in order to satisfy the timing requirements of the Flash memory.

The length of the ring depends on the process used for manufacturing the chip. A nominal operating frequency range of 10 MHz is sufficient. This clock should contain a small amount of randomization to prevent attacks where light emissions from switching events are captured.

Note that this is different to the input SClk which is the serial clock for external communication.

The ring oscillator is covered by both Tamper Detection and Prevention lines so that if an attacker attempts to tamper with the unit, the chip will either RESET or erase all secret information.

FPGA Note: the FPGA does not have an internal ring oscillator. An additional pin (SysClk) is used instead. This is replaced by an internal ring oscillator in the final ASIC.

11.1.2 Voltage Reference

The voltage reference block maintains an output which is substantially independent of process, supply voltage and temperature. It provides a reference voltage which is used by the PDU and a reference current to stabilise the ring oscillator. It may also be used as part of the temperature based clock filter described in Section 10.3.3 on page 961.

11.1.3 OverUnder Power Detection Unit

The OverUnder Power Detection Unit (PDU) is the same as that described in Section 10.3.9 on page 965.

The Under Voltage Detection Unit provides the signal PwrFailing which, if asserted, indicates that the power supply may be turning off. This signal is used to rapidly terminate any Flash write that may be in progress to avoid accidentally writing to an indeterminate memory location.

Note that the PDU triggers the RESET Tamper Detection Line only. It does not trigger the Erase Tamper Detection Line.

The PDU can be implemented with regular CMOS, since the key does not pass through this unit. It does not have to be implemented with non-flashing CMOS.

The PDU is covered by both Tamper Detection and Prevention lines so that if an attacker attempts to tamper with the unit, the chip will either RESET or erase all secret information.

11.1.4 Power-On Reset and Tamper Detect Unit

The Power-on Reset unit (POR) detects a power-on condition and generates the PORstL signal that is fed to all the validation units, including the two inside the Tamper Detect Unit (TDU).

All other logic is connected to RstL, which is the PORstL gated by the VAL unit attached to the Reset tamper detection lines (see Section 10.3.5 on page 962) within the TDU. Therefore, if the Reset tamper line is asserted, the validation will drive RstL low, and can only be cleared by a power-down. If the tamper line is not asserted, then RstL=PORstL.

The TDU contains a second VAL unit attached to the Erase tamper detection lines (see Section 10.3.5 on page 962) within the TDU. It produces a TamperEraseOK signal that is output to the MIU (1=the tamper lines are all OK, 0=force an erasure of Flash).

11.1.5 Noise Generator

The Noise Generator (NG) is the same as that described in Section 10.3.4 on page 961. It is based on a 64-bit maximal period LFSR loaded with a set non-zero bit pattern on RESET.

The NG must be protected by both Tamper Detection and Prevention lines so that if an attacker attempts to tamper with the unit, the chip will either RESET or erase all secret information.

In addition, the bits in the LFSR must be validated to ensure they have not been tampered with (i.e. a parity check). If the parity check fails, the Erase Tamper Detection Line is triggered.

Finally, all 64 bits of the NG are ORed into a single bit. If this bit is 0, the Erase Tamper Detection Line is triggered. This is because 0 is an invalid state for an LFSR.

11.2 Trim Unit

The 8-bit Trim register within the Trim Unit has a reset value of 0x00(to enable the flash reads to succeed even in the fastest process corners), and is written to either by the PMU during Trim Mode or by the CPU in Active Mode. Note that the CPU is only able to write once to the Trim register between power-on-reset due to the TrimDone flag which provides overloading of LocalidWE.

The reset value of Trim (0) means that the chip has a nominal frequency of 2.7 MHz-10 MHz. The upper of the range is when we cannot trim it lower than this (or we could allow some spread on the acceptable trimmed frequency but this will reduce our tolerance to ageing, voltage and temperature which is the range 7 MHz to 14 MHz). The 2.7 MHz value is determined by a chip whose oscillator runs at 10 MHz when the trim register is set to its maximum value, so then it must run at 2.7 MHz when trim=0. This is based on the non-linear frequency-current characteristic of the oscillator.

Chips found outside of these limits will be rejected.

The frequency of the ring oscillator is measured by counting cycles⁶, in the PMU, over the byte period of the serial interface. The frequency of the serial clock, SClk, and therefore the byte period will be accurately controlled during the measurement. The cycle count (Fineas) at the end of the period is read over the serial bus and the Trim register updated (Trimval) from its power on default (POD) value. The steps are shown in FIG. 391. Multiple measure—read—trim cycles are possible to improve the accuracy of the trim procedure. 6Note that the PMU counts using 12-bits, saturates at 0xFFF, and returns the cycle count divided by 2 as an 8-bit value. This means that multiple measure-read-trim cycles may be necessary to resolve any amibguity. In any case, multiple cycles are necessary to test the correctness of the trim circuitry during manufacture test.

A single byte for both Fineas and Trimval provide sufficient accuracy for measurement and trimming of the frequency. If the bus operates at 400 kHz, a byte (8 bits) can be sent in 20 μs. By dividing the maximum oscillator frequency, expected to be 20 MHz, by 2 results in a cycle count of 200 and 50 for the minimum frequency of 5 MHz resulting in a worst case accuracy of 2%.

FIG. 392 shows a block diagram of the Trim Unit:

The 8-bit Trim value is used in the analog Trim Block to adjust the frequency of the ring oscillator by controlling its bias current. The two lsbs are used as a voltage trim, and the 6 msbs are used as a frequency trim.

The analog Trim Clock circuit also contains a Temperature filter as described in Section 10.3.3 on page 961.

11.310 Unit

The QA Chip acts as a slave device, accepting serial data from an external master via the IO Unit (IOU). Although the IOU actually transmits data over a 1-bit line, the data is always transmitted and received in 1-byte chunks.

The IOU receives commands from the master to place it in a specific operating mode, which is one of:

-   -   Idle Mode: is the startup mode for the IOU if the fuse has not         yet been blown. Idle Mode is the mode where the QA Chip is         waiting for the next command from the master. Input signals from         the CPU are ignored.     -   Program Mode: is where the QA Chip erases all currently stored         data in the Flash memory (program and secret key information)         and then allows new data to be written to the Flash. The IOU         stays in Program Mode until told to enter another mode.     -   Active Mode: is the startup mode for the IOU if the fuse has         been blown (the program is safe to run). Active Mode is where         the QA Chip allows the program code to be executed to process         the master's specific command. The IOU returns to Idle Mode         automatically when the command has been processed, or if the         time taken between consuming input bytes (while the master is         writing the data) or generating output bytes (while the master         is reading the results) is too great.     -   Trim Mode: is where the QA Chip allows the generation and         setting of a trim value to be used on the internal ring         oscillator clock value. This must be done for safety reasons         before a program can be stored in the Flash memory.

See Section 12 on page 970 for detailed information about the IOU.

11.4 Central Processing Unit

The Central Processing Unit (CPU) block provides the majority of the circuitry of the 4-bit microprocessor. FIG. 393 shows a high level view of the block.

11.5 Memory Interface Unit

The Memory Interface Unit (MIU) provides the interface to flash and RAM. The MIU contains a Program Mode Unit that allows flash memory to be loaded via the IOU, a Memory Request Unit that maps 8-bit and 32-bit requests into multiple byte based requests, and a Memory Access Unit that generates read/write strobes for individual accesses to the memory.

FIG. 394 shows a high level view of the MIU block.

11.6 Memory Components

The Memory Components block isolates the memory implementation from the rest of the QA Chip. The entire contents of the Memory Components block must be protected from tampering. Therefore the logic must be covered by both Tamper Detection Lines. This is to ensure that program code, keys, and intermediate data values cannot be changed by an attacker. The 8-bit wide RAM also needs to be parity-checked.

FIG. 395 shows a high level view of the Memory Components block. It consists of 8 KBytes of flash memory and 3072 bits of parity checked RAM.

11.6.1 RAM

The RAM block is shown here as a simple 96×32-bit RAM (plus parity included for verification).

The parity bit is generated during the write.

The RAM is in an unknown state after RESET, so program code cannot rely on RAM being 0 at startup.

The initial version of the ASIC has the RAM implemented by Artisan component RA1SH (96×32-bit RAM without parity). Note that the RAMOutEn port is active low i.e. when 0, the RAM is enabled, and when 1, the RAM is disabled.

11.6.2 Flash Memory

A single Flash memory block is used to hold all non-volatile data. This includes program code and variables. The Flash memory block is implemented by TSMC component SFC0008_(—)08B9_HE [4], which has the following characteristics:

-   -   8 K×8-bit main memory, plus 128×8-bit information memory     -   512 byte page erase     -   Endurance of 20,000 cycles (min)     -   Greater than 100 years data retention at room temperature     -   Access time: 20 ns (max)     -   Byte write time: 20 μs (min)     -   Page erase time: 20 ms (min)     -   Device erase time: 200 ms (min)     -   Area of 0.494 mm² (724.66 μm×682.05 μm)

The FlashCtrl line are the various inputs on the SFC0008_(—)08B9_HE required to read and write bytes, erase pages and erase the device. A total of 9 bits are required (see [4] for more information).

Flash values are unchanged by a RESET. After manufacture, the Flash contents must be considered to be garbage. After an erasure, the Flash contents in the SFC0008_(—)08B9_HE is all 1 s.

11.6.3 VAL Blocks

The two VAL units are validation units connected to the Tamper Prevention and Detection circuitry (described in Section 10.3.5 on page 962), each with an OK bit. The OK bit is set to 1 on PORstL, and ORed with the ChipOK values from both Tamper Detection Lines each cycle. The OK bit is ANDed with each data bit that passes through the unit.

In the case of VAL₁, the effective byte output from the flash will always be 0 if the chip has been tampered with. This will cause shadow tests to fail, program code will not execute, and the chip will hang.

In the case of VAL₂, the effective byte from RAM will always be 0 if the chip has been tampered with, thus resulting in no temporary storage for use by an attacker.

12 IO Unit

The I/O Unit (IOU) is responsible for providing the physical implementation of the logical interface described in Section 5.1 on page 933, moving between the various modes (Idle, Program, Trim and Active) according to commands sent by the master.

The IOU therefore contains the circuitry for communicating externally with the external world via the SClk and SDa pins. The IOU sends and receives data in 8-bit chunks. Data is sent serially, most significant bit (bit 7) first through to least significant bit (bit 0) last. When a master sends a command to an QA Chip, the command commences with a single byte containing an id in bits 7-1, and a read/write sense in bit 0, as shown in FIG. 396.

The IOU recognizes a global id of 0x00 and a local id of LocalId (set after the CPU has executed program code at reset or due to a global id/ActiveMode command on the serial bus). Subsequent bytes contain modal information in the case of global ID, and command/data bytes in the case of a match with the local id.

If the master sends data too fast, then the IOU will miss data, since the IOU never holds the bus.

The meaning of too fast depends on what is running. In Program Mode, the master must send data a little slower than the time it takes to write the byte to flash (actually written as 2×8-bit writes, or 40 μs). In ActiveMode, the master is permitted to send and request data at rates up to 500 KHz.

None of the latches in the IOU need to be parity checked since there is no advantage for an attacker to destroy or modify them.

The IOU outputs 0s and inputs 0s if either of the Tamper Detection Lines is broken. This will only come into effect if an attacker has disabled the RESET and/or erase circuitry, since breaking either Tamper Detection Lines should result in a RESET or the erasure of all Flash memory.

The IOU's InByte, InByteValID, OutByte, and OutByteValid registers are used for communication between the master and the QA Chip. InByte and InByteValid provide the means for clients to pass commands and data to the QA Chip. OutByte and OutByteValid provide the means for the master to read data from the QA Chip.

-   -   Reads from InByte should wait until InByteValid is set.         InByteValid will remain clear until the master has written the         next input byte to the QA Chip. When the IOU is told (by the FEU         or MU) that InByte has been read, the IOU clears the InByteValid         bit to allow the next byte to be read from the client.     -   Writes to OutByte should wait until OutByteValid is clear.         Writing OutByte sets the OutByteValid bit to signify that data         is available to be transmitted to the master. OutByteValid will         then remain set until the master has read the data from OutByte.         If the master requests a byte but OutByteValid is clear, the IOU         sends a NAck to indicate the data is not yet ready.

When the chip is reset via RstL, the IOU enters ActiveMode to allow the PMU to run to load the fuse. Once the fuse has been loaded (when MlUAvail transitions from 0 to 1) the IOU checks to see if the program is known to be safe. If it is not safe, the IOU reverts to IdleMode. If it is safe (FuseBlown=1), the IOU stays in ActiveMode to allow the program to load up the localId and do any other reset initialization, and will not process any further serial commands until the CPU has written a byte to the OutByte register (which may be read or not at the discretion of the master using a localId read). In both cases the master is then able to send commands to the QA Chip as described in Section 5.1 on page 933.

FIG. 397 shows a block diagram of the IOU.

With regards to InByteValid inputs, set has priority over reset, although both set and reset in correct operation should never be asserted at the same time. With regards to IOSetInByte and IOLoadInByte, if IOSetInByte is asserted, it will set InByte to be 0xFF regardless of the setting of IOLoadInByte.

The two VAL units are validation units connected to the Tamper Prevention and Detection circuitry (described in Section 10.3.5 of the Architecture Overview chapter), each with an OK bit. The OK bit is set to 1 on PORstL, and ORed with the ChipOK values from both Tamper Detection Lines each cycle. The OK bit is ANDed with each data bit that passes through the unit.

In the case of VAL₁, the effective byte output from the chip will always be 0 if the chip has been tampered with. Thus no useful output can be generated by an attacker. In the case of VAL2, the effective byte input to the chip will always be 0 if the chip has been tampered with. Thus no useful input can be chosen by an attacker.

There is no need to verify the registers in the IOU since an attacker does not gain anything by destroying or modifying them.

The current mode of the IOU is output as a 2-bit IOMode to allow the other units within the QA Chip to take correct action. IOMode is defined as shown in Table 372:

TABLE 372 IOMode values Value Interpretation 00 Idle Mode 01 Program Mode 10 Active Mode 11 Trim Mode

The Logic blocks generate a 1 if the current IOMode is in Program Mode, Active Mode or Trim Mode respectively. The logic blocks are:

Logic₁ IOMode = 01 (Program) Logic₂ IOMode = 10 (Active) Logic₃ IOMode = 11 (Trim) 12.1 State Machine

There are two state machines in the IOU running in parallel. The first is a byte-oriented state machine, the second is a bit-oriented state machine. The byte-oriented state machine keeps track of the operating mode of the QA Chip while the bit-oriented state machine keeps track of the low-level bit Rx/Tx protocol.

The SDa and SClk lines are connected to the respective pads on the QA Chip. The IOU passes each of the signals from the pads through 2 D-types to compensate for metastability on input, and then a further latch and comparitor to ensure that signals are only used if stable for 2 consecutive internal clock cycles. The circuit is shown in Section 12.1.1 below.

12.1.1 Start/Stop Control Signals

The StartDetected and StopDetected control signals are generated based upon monitoring SDa synchronized to SClk. The StartDetected condition is asserted on the falling edge of SDa synchronized to SClk, and the StopDetected condition is asserted on the rising edge of SDa synchronized to SClk.

In addition we generate feSClk which is asserted on the falling edge of SClk, and reSClk which is asserted on the rising edge of SClk. Finally, feSclkPrev is the value of feSClk delayed by a single cycle. FIG. 398 shows the relationship of inputs and the generation of SDaReg, reSClk, feSClk, feSclkPrev, StartDetected and StopDetected.

The SDaRegSelect logic compensates for the 2:1 variation in clock frequency. It uses the length of the high period of the SClk (from the saturating counter) to select between sda5, sda6 and sda7 as the valid data from 300 ns before the falling edge of SClk as follows.

The minimum time for the high period of SClk is 600 ns. If the counter <=4 (i.e. 5 or fewer cycles with SClk=1) then SDaReg output=sda5 (sample point is equidistant from rising and falling edges). If the counter=5 or 6 (i.e. 6 or 7 samples where SClk=1), then SDaReg output=sda6. If the counter=7 (the counter saturates when there are 8 samples of SClk=1), then SDaReg output=sda7. This is shown in pseducode below:

If ((counter₂ = 0)

(counter = 4)) SDaReg = sda5 ElseIf (counter = 7) SDaReg = sda7 Else SDaReg = sda6 EndIf

The counter also provides a means of enabling start and stop detection. There is a minimum of a 600 ns setup and 600 ns hold time for start and stop conditions. At 14 MHz this means samples 4 and 5 after the rising edge (sample 1 is considered to be the first sample where SClk=1) could potentially include a valid start or stop condition. At 7 MHz samples 4 and 5 represent 284 and 355 ns respectively, although this is after the rising edge of SClk, which itself is 100 ns after the setup of data (i.e. 384 and 455 ns respectively and therefore safe for sampling). Thus the data will be stable (although not a start or stop). Since we detect stops and starts using sda5 and sda6, we can only validly detect starts and stops 6 cycles after a rising edge, and we need to not-detect starts and stops 4 cycles before the falling edge. We therefore only detect starts and stops when the counter is >=6 (i.e. when sclk3 and scik2 are 0 and 1 respectively, sda2 holds sample 1 coincident with the rising edge, sdal holds sample 2, sda0 holds sample 3, we load the counter with 0 and sample SDa to obtain the new sda0 which will hold sample 4 at the end of the cycle. Thus while the counter is incrementing from 0 to 1, sda0 will hold sample 4. Therefore sample 4 will be in sda6 when the counter is 6.

12.1.2 Control of SDa and SClk Pins

The SClk line is always driven by the master. The SDa line is driven low whenever we want to transmit an ACK (SDa is active low) or a 0-bit from OutByte. The generation of the SDa pin is shown in the following pseudocode:

TxAck = (bitSM_state = ack)

((byteSM_state = doWrite)

(((byteSM_state = getGlobalCmd)

(byteSM_state = checkId))

AckCmd)) TxBit

(byteSM_state = doRead)

(bitSM_state = xferBit)

OutByte

_(bitCount) SDa =

(TxAck

TxBit) # only drive the line when we are xmitting a 0

The slew rate of the SDa line should be restricted to minimise ground bounce. The pad must guarantee a fall time >20 ns. The rise time will be controlled by the external pull up resistor and bus capacitance.

12.1.3 Bit-Oriented State Machine

The bit-oriented state machine keeps track of the general flow of serial transmission including start/data/ack/stop as shown in the following pseudocode:

idle EndByte = FALSE EndAck = FALSE If (StartDetected) state

starting Else state

idle EndIf starting EndByte = FALSE EndAck = FALSE NAck

0 If (StopDetected) state

idle ElseIf (feSClkPrev) bitCount

0 state

xferBit Else state

starting# includes StartDetected EndIf xferBit EndAck = FALSE EndByte = (feSclkPrev

(bitCount = 0)) # after feSclk bitCount must be 1..8 If (feSClk) shiftLeft[ioByte, SDaReg] # capture the bit in the ioByte shift register bitCount

bitCount + 1 # modulo count due to 3 bit bitCount EndIf If (StopDetected) state

idle ElseIf (StartDetected) state

starting ElseIf (EndByte) state

ack Else state

xferBit EndIf ack EndByte = FALSE EndAck = feSclkPrev If (StopDetected) state

idle ElseIf (StartDetected) state

starting ElseIf (EndAck) state

xferBit # bitCount is already 0 Else If (feSClk) NAck

SDaReg # active low, so 0 = ACK, 1 = NACK EndIf state

ack EndIf 12.1.4 Byte-Oriented State Machine

The following pseudocode illustrates the general startup state of the IOU and the receipt of a transmission from the master.

rstL # setup state of registers on reset IOMode

ActiveMode # to force the fuse to be loaded OutByteValid

0 OutByte

0 InByteValid

1 # required InByte

0xFF # byte = FF = the ‘reset’ command localId

0  # loads localId with the globalId so no localId exists state

wait4fuse wait4fuse If (MIUAvail) If (FuseBlown) # this must be done same cycle as seeing MIUAvail go high state

wait4cpu Else IOMode

IdleMode # CPU will now require an external ActiveMode to start state

idle Else state

wait4fuse EndIf wait4cpu If (CPUOutByteWE) # wait for CPU reset activities to finish state

idle # note: we're still in ActiveMode Else state

wait4cpu EndIf idle If (StartDetected) state

checkId Else state

idle EndIf

The first byte received must be checked to ensure it is meant for everyone (globaild of 0) or specifically for us (localid matches). We only send an ACK to a read when there is data available to send. In addition, writes to the general call address (0) are always ACKed, but reads from the general call address are only ACKed before the fuse has been blown.

checkId isWrite = (ioByte₀ = 0) isRead = (ioByte₀ = 1) isGlobal = (ioByte⁷⁻¹ = 0) globalW = isGlobal

isWrite localW = (ioByte⁷⁻¹ = localID)

isWrite

isGlobal localR = (ioByte⁷⁻¹ = localID)

isRead

(

GlobalW

FuseBlown) If (StopDetected) state

idle ElseIf (EndByte) AckCmd_in = (globalW

localW)

(localR

OutByteValid) AckCmd

AckCmd_in If (localW) IOMode

IdleMode # jic - any output was pending IOOutByteUsed = 1 IOClearInByte = 1 # ensure there is nothing hanging around from before EndIf ElseIf (EndAck) If (globalW) # globalW and localW are mutually exclusive state

getGlobalCmd ElseIf (localW) IOMode

ActiveMode IOLoadInByte = 1 # will set inByte to localW (lsb will be 0) state

doWrite ElseIf (localR

IOMode₁

AckCmd) # Active mode (or Trim when fuse intact) state

doRead Else state

idle # ignore reads unless first in active or trim mode EndIf Else state

checkId EndIf

With a new global command the IOU waits for the mode byte (see Table page6 on page 934) to determine the new operating mode:

getGlobalCmd wantProg = ((ioByte = ProgramModeId)

FuseBlown) wantTrim = ((ioByte = TrimModeId)

FuseBlown) wantActive = (ioByte = ActiveModeId) If (StopDetected) state

idle ElseIf (StartDetected) state

checkId ElseIf (EndByte) AckCmd_in = wantActive

wantProg

wantTrim # only ACK cmds we can do AckCmd

AckCmd_in If (AckCmd_in) IOMode

IdleMode # jic - any output was pending IOOutByteUsed = 1 IOClearInByte = 1 # ensure there is nothing hanging around from before EndIf ElseIf (EndAck) If (wantProg) IOMode

ProgramMode # don't load inByte (we only want the data) state

doWrite ElseIf (wantTrim) IOMode

TrimMode # don't load InByte (we only want the next byte) state

doWrite ElseIf (wantActive) # must be Active IOMode

ActiveMode IOSetInByte = 1 # 0 for all other cases & states. 1 = sets inByte to 0xFF IOLoadInByte = 1 # sets InByteValid (InByte is set to 0xFF (‘reset’ cmd)) state

wait4cpu# don't do anything til the cpu has completed this task Else state

idle # unknown id, so ignore remainder EndIf Else state

getGlobalCmd EndIf

When the master writes bytes to the QA Chip (e.g. parameters for a command), the program must consume the byte fast enough (i.e. during the sending of the ACK) or subsequent bits may be lost.

The process of receiving bytes is shown in the following pseudocode:

doWrite If (StopDetected) state

idle # stay in whatever IOMode we were in ElseIf (StartDetected) state

checkId Else If (EndByte) IOLoadInByte =

InByteValid EndIf If (EndByte

InByteValid) # will only be when master sends data too quickly state

idle # ACK will not be sent when in idle state Else state

doWrite # ACK will be sent automatically after byte is Rxed EndIf EndIf

When the master wants to read, the IOU sends one byte at a time as requested. The process is shown in the following pseudocode:

doRead If (StopDetected) state

idle ElseIf (StartDetected) state

checkId ElseIf (EndAck) If (NAck

OutByteValid) state

idle Else state

doRead EndIf Else If (EndByte) IOOutByteUsed = 1 EndIf state

doRead EndIf 13 Fetch and Execute Unit 13.1 Introduction

The QA Chip does not require the high speeds and throughput of a general purpose CPU. It must operate fast enough to perform the authentication protocols, but not faster. Rather than have specialized circuitry for optimizing branch control or executing opcodes while fetching the next one (and all the complexity associated with that), the state machine adopts a simplistic view of the world. This helps to minimize design time as well as reducing the possibility of error in implementation.

The FEU is responsible for generating the operating cycles of the CPU, stalling appropriately during long command operations due to memory latency.

When a new transaction begins, the FEU will generate a JPZ (jump to zero) instruction.

The general operation of the FEU is to generate sets of cycles:

-   -   Cycle 0: fetch cycles. This is where the opcode is fetched from         the program memory, and the effective address from the fetched         opcode is generated. The Fetch output flag is set during the         final cycle 0 (i.e. when the opcode is finally valID).     -   Cycle 1: execute cycle. This is where the operand is         (potentially) looked up via the generated effective address         (from Cycle 0) and the operation itself is executed. The Exec         output flag is set during the final cycle 1 (i.e. when the         operand is finally valID).

Under normal conditions, the state machine generates multiple Cycle=0 followed by multiple Cycle=1. This is because the program is stored in flash memory, and may take multiple cycles to read. In addition, writes to and erasures of flash memory take differing numbers of cycles to perform. The FEU will stall, generating multiple instances of the same Cycle value with Fetch and Exec both 0 until the input MIURdy=1, whereupon a Fetch or Exec pulse will be generated in that same cycle.

There are also two cases for stalling due to serial I/O operations:

-   -   The opcode is ROR OutByte, and OutByteValid=1. This means that         the current operation requires outputting a byte to the master,         but the master hasn't read the last byte yet.     -   The operation is ROR InByte, and InByteValid=0. This means that         the current operation requires reading a byte from the master,         but the master hasn't supplied the byte yet.

In both these cases, the FEU must stall until the stalling condition has finished.

Finally, the FEU must stop executing code if the IOU exits Active Mode.

The local Cmd opcode/operand latch needs to be parity-checked. The logic and registers contained in the FEU must be covered by both Tamper Detection Lines. This is to ensure that the instructions to be executed are not changed by an attacker.

13.2 State Machine

The Fetch and Execute Unit (FEU) is combinatorial logic with the following registers:

TABLE 373 FEU Registers Name #bits Description Output registers (visible outside the FEU) Cycle 1 0 if the FEU is currently fetching an opcode, 1 if the FEU is currently executing the opcode. NewMemTrans 1 Is asserted during the start of a potential new memory access. 0 = this is not the first cycle of a set of Cycle 0 or Cycle 1 1 = this is the first cycle of a set of Cycle 0 or Cycle 1 (previous cycle must have been a Fetch or an Exec). Go 1 1 if the FEU is currently fetching and executing program code (i.e. a program is currently running), 0 if it is not. Local registers (not visible outside the FEU) CurrCmd 8 + p Holds the currently executing instruction (parity checked). PendingKill 1 The currently executing program is waiting to be halted (waiting due to memory access) PendingStart 1 A new transaction is waiting to be started (waiting due to memory access or an existing transaction not yet stopped) WasIdle 1 The previous cycle had an IOMode of IdleMode.

In addition, the following externally visible outputs are generated asynchronously:

TABLE 374 Externally visible asynchronous FEU outputs Name #bits Description Fetch 1 1 if the FEU is performing the final cycle of a fetch (i.e. Cycle will also be 0). It is set when the NextCmd output is valid. The local Cmd register is latched during the Fetch cycle with either the incoming MIU8Data or an FEU-generated command. Exec 1 1 if the FEU is performing the final cycle of an execute (i.e. Cycle will also be 1). It is set when the data required by the opcode from the MIU is valid. Other units can execute the Cmd and latch data from the MIU (e.g. from MIUData) during the Exec cycle. Cmd 8 When Cycle = 0, this holds the next instruction to be executed (during the next Cycle = 1). Is generated based on incoming MIU8Data or substituted FEU command (e.g. JSR 0). When Cycle = 1, this holds the current instruction being executed (based on theCmd).

The Cycle and currCmd registers are not used directly. Instead, their outputs are passed through a VAL unit before use. The VAL units are designed to validate the data that passes through them. Each contains an OK bit connected to both Tamper Prevention and Detection Lines. The OK bit is set to 1 on PORstL, and ORed with the ChipOK values from both Tamper Detection Lines each cycle. The OK bit is ANDed with each data bit that passes through the unit.

In the case of VAL₁, the effective Cycle will always be 0 if the chip has been tampered with. Thus no program code will execute.

In the case of VAL₂, the effective 8-bit currCmd value will always be 0 if the chip has been tampered with. Multiple 0s will be interpreted as the JSR 0 instruction, and this will effectively hang the CPU. VAL₂ also performs a parity check on the bits from currCmd to ensure that currCmd has not been tampered with. If the parity check fails, the Erase Tamper Detection Line is triggered. For more information on Tamper Prevention and Detection circuitry, see Section 10.3.5 on page 962.

13.2.1 Pseudocode

reset conditions: Fetch = 0 Exec = 0 Cycle

0 currCmd

0 Go

0 pendingKill

0 pendingStart

0 newMemTrans

0 wasIdle

1 # required to detect if IOU starts in a non-idle state

The cycle by cycle combinatorial logic behaviour is shown in the following pseudocode:

isActive = (IOMode = ActiveMode) wasIdle

(IOMode = IdleMode) wantToStart = (pendingStart

wasIdle)

isActive newTrans = wantToStart

Go

MIUAvail pendingStart

wantToStart

newTrans killTrans = Go

(

isActive

pendingKill) Fetch = newTrans

(Go

Cycle

MIURdy

killTrans) inDelay = (currCmd = ROR InByte)

InByteValid outDelay = (currCmd = ROR OutByte)

OutByteValid ioDelay = inDelay

outDelay Exec = Go

Cycle

MIURdy

ioDelay If (Cycle) Cmd = currCmd ElseIf (newTrans) Cmd = JPZ # jump to 0 Else Cmd = MIU8Data EndIf resetGo = (MIURdy

killTrans)

(Fetch

(Cmd = HALT)) pendingKill

killTrans

resetGo changeCycle = Fetch

Exec # will only be 1 when Go = 1 Cycle

newTrans

((Cycle ⊕ changeCycle)

resetGo) newMemTrans

newTrans

(changeCycle

resetGo) If (Fetch) currCmd

Cmd EndIf If (resetGo) Go

0 ElseIf (newTrans) Go

1 EndIf 14 ALU

The Arithmetic Logic Unit (ALU) contains a 32-bit Acc (Accumulator) register as well as the circuitry for simple arithmetic and logical operations.

The logic and registers contained in the ALU must be covered by both Tamper Detection Lines. This is to ensure that keys and intermediate calculation values cannot be changed by an attacker. In addition, the Accumulator must be parity-checked.

A 1-bit Z signal represents the state of zero-ness of the Accumulator. The Accumulator is cleared to 0 upon a RstL, and the Z signal is set to 1. The Accumulator is updated for any of the commands: AND, OR, XOR, ADD, ROR, and RIA, and the Z signal is updated whenever the Accumulator is updated. Note that the Z signal is actually implemented as a nonZ register whose output is passed through an inverter and used as Z.

Each arithmetic and logical block operates on two 32-bit inputs: the current value of the Accumulator, and the current 32-bit output of the DataSel block (either the 32 bit value from MIUData or an immediate value). The AND, OR, XOR and ADD blocks perform the standard 32-bit operations. The remaining blocks are outlined below.

FIG. 399 shows a block diagram of the ALU:

The Accumulator is updated for all instructions where the high bit of the opcode is set:

Logic₁ Exec

Cmd₇

Since the WriteEnables of Acc and nonZ takes Cmd7 and Exec into account (due to Logic₁), these two bits are not required by the multiplexor MX₁ in order to select the output. The output selection for MX₁ only requires bits 6-3 of the Cmd and is therefore simpler as a result (as shown in Table 375).

TABLE 375 Selection for multiplexor MX₁ Output Cmd⁶⁻³ MX₁ immOut 011x

1110 (LD) rorOut 100x

1111 (RIA, ROR) from XOR 001x

1100 (XOR) from ADD 010x

1101 (ADD) from AND 0000

1010 (AND) from OR 0001

1011 (OR)

The two VAL units are validation units connected to the Tamper Prevention and Detection circuitry (described in Section 10.3.5 on page 962), each with an OK bit. The OK bit is set to 1 on PORstL, and ORed with the ChipOK values from both Tamper Detection Lines each cycle. The OK bit is ANDed with each data bit that passes through the unit.

In the case of VAL₁, the effective bit output from the Accumulator will always be 0 if the chip has been tampered with. This prevents an attacker from processing anything involving the Accumulator. VAL₁ also performs a parity check on the Accumulator, setting the Erase Tamper Detection Line if the check fails.

In the case of VAL₂, the effective Z status of the Accumulator will always be true if the chip has been tampered with. Thus no looping constructs can be created by an attacker.

14.1 DataSel Block

The DataSel block is designed to implement the selection between the MIU32Data and the immediate addressing mode for logical commands.

Immediate addressing relies on 3 bits of operand, plus an optional 8 bits at PC+1 to determine an 8-bit base value. Bits 0 to 1 determine whether the base value comes from the opcode byte itself, or from PC+1, as shown in Table 376.

TABLE 376 Selection for base value in immediate mode Cmd¹⁻⁰ Base value 00 00000000 01 00000001 10 From PC + 1 (i.e. MIUData³¹⁻²⁴) 11 11111111

The base value is computed by using CMD₀ as bit 0, and copying CMD₁ into the upper 7 bits.

The 8-bit base value forms the lower 8 bits of output. These 8 bits are also ANDed with the sense of whether the data is replicated in the upper bits or not (i.e. CMD₂). The resultant bits are copied in 3 times to form the upper 24 bits of the output.

FIG. 400 shows a block diagram of the ALU's DataSel block:

14.2 ROR Block

The ROR block implements the ROR and RIA functionality of the ALU.

A 1-bit register named RTMP is contained within the ROR unit. RTMP is cleared to 0 on a RstL, and set during the ROR RB and ROR XRB commands. The RTMP register allows implementation of Linear Feedback Shift Registers with any tap configuration.

FIG. 401 shows a block diagram of the ALU's ROR block:

The ROR n, blocks are shown for clarity, but in fact would be hardwired into multiplexor MX₃, since each block is simply a rewiring of the 32-bits, rotated right n bits.

Logic₁ is used to provide the WriteEnable signal to RTMP. The RTMP register should only be written to during ROR RB and ROR XRB commands. The combinatorial logic block is:

Logic₁ Exec

(Cmd⁷⁻⁴ = ROR)

(Cmd³⁻¹ = 000)

Multiplexor MX₁ performs the task of selecting the 6-bit value from Cn instead of bits 13-8 (6 bits) from Acc (the selection is based on the value of Logic₂). Bit 5 is required to distinguish ROR from RIA.

Logic₂ Cmd⁵⁻² = 0x10

TABLE 377 Selection for multiplexor MX₁ Output Logic₂ MX₁ Cn 1 Acc¹³⁻⁸ 0

Multiplexor MX₂ performs the task of selecting the 8-bit value from InByte instead of the lower 8 bits from the ANDed Acc based on the CMD.

TABLE 378 Selection for multiplexor MX₂ Output Cmd⁴⁻⁰ MX₂ InByte 0x110 Acc⁷⁻⁰

(0x110)

Multiplexor MX₃ does the final rotating of the 32-bit value. The bit patterns of the CMD operand are taken advantage of:

TABLE 379 Selection for multiplexor MX₃ Output Cmd³⁻⁰ Comments MX₃ ROR 1 00xx RB, XRB, WriteMask, 1 ROR 3 010x  3 ROR 31 0110 31 ROR 24 0111 24 ROR 8 1xxx RIA, InByte, 8, OutByte, C1, C2, ID 14.3 IO Block

The IO block within the ALU implements the logic for communicating with the IOU during instructions that involve the Accumulator. This includes generating appropriate control signals and for generating the correct data for sending during writes to the IOU's OutByte and LocalId registers.

FIG. 402 shows a block diagram of the IO block:

Logic₁ is used to provide the LocalIdWE signal to the IOU. The localId register should only be written to during the ROR ID command. Only the lower 7 bits of the Accumulator are written to the localId register.

Logic₂ is used to provide the ALUOutByteWE signal to the IOU. The OutByte register should only be written to during the ROR OutByte command. Only the lower 8 bits of the Accumulator are written to the OutByte register.

In both cases we output the lower 8 bits of the Accumulator. The ALUIOData value is ANDed with the output of Logic₂ to ensure that ALUIOData is only valid when it is safe to do so (thus the IOU logic never sees the key passing by in ALUIOData). The combinatorial logic blocks are:

Logic₁ Exec

(Cmd⁷⁻⁰ = ROR ID) Logic₂ Exec

(Cmd⁷⁻⁰ = ROR OutByte)

Logic₃ is used to provide the ALUInByteUsed signal to the IOU. The InByte is only used during the ROR InByte command. The combinatorial logic is:

Logic₃ Exec

(Cmd⁷⁻⁰ = ROR InByte) 15 Program Counter Unit

The Program Counter Unit (PCU) includes the 12 bit PC (Program Counter), as well as logic for branching and subroutine control.

The PCU latches need to be parity-checked. In addition, the logic and registers contained in the PCU must be covered by both Tamper Detection Lines to ensure that the PC cannot be changed by an attacker.

The PC is implemented as a 12 entry by 12-bit PCA (PC Array), indexed by a 4-bit SP (Stack Pointer) register. The PC, PCRamSel and SP registers are all cleared to 0 on a RstL, and updated during the flow of program control according to the opcodes.

The current value for the PC is normally updated during the Execute cycle according to the command being executed. However it is also incremented by 1 during the Fetch cycle for two byte instructions such as JMP, JSR, DBR, TBR, and instructions that require an additional byte for immediate addressing. The mechanism for calculating the new PC value depends upon the opcode being processed.

FIG. 403 shows a block diagram of the PCU:

The ADD block is a simple adder modulo 2¹² with two inputs: an unsigned 12 bit number and an 8-bit signed number (high bit=sign). The signed input is either a constant of 0x01, or an 8-bit offset (the 8 bits from the MIU).

The “+1.” block takes a 4-bit input and increments it by 1 (modulo 12). The “−1.” block takes a 4-bit input and decrements it by 1 (modulo 12).

Table 380 lists the different forms of PC control:

TABLE 381 Different forms of PC control during the Exec cycle Command Action JMP The PC is loaded with the current 12-bit value as passed in from the MIU. JPI The PC is loaded with the current 12-bit value as passed in from the Acc. PCRamSel is loaded with the value from bit 15 of the Acc. JPZ The PC is loaded with 0. PCRamSel is loaded with 0 (program in flash) JSZ Save old value of PC onto stack for later. The PC is loaded with 0. PCRamSel is loaded with 0 (program in flash). JSR, JSI Save old value of PC onto stack for later. The PC is loaded with the current 12-bit value as passed in from either the MIU or the Acc. With JSI, PCRamSel is loaded from the value in bit 15 of the Accumulator. RTS Pop old value of PC from stack and increment by 1 to get new PC. TBR If the Z flag matches the TBR test, add 8-bit signed number (MIU8Data) to current PC. Otherwise increment current PC by 1. DBR If the CZ flag is set, add 8-bit signed offset (MIU8Data) to current PC. Otherwise increment current PC by 1. All others Increment current PC by 1

The updating of PCRamSel only occurs during JPI, JSI, JPZ and JSZ instructions, detected via Logic₀.

The same action for the Exec takes place for JMP, JSR, JPI, JSI, JPZ and JSZ, so we specifically detect that case in Logic₁. In the same way, we test for the RTS case in Logic₂.

Logic₀ Cmd⁷⁻¹ = 011x001 Logic₁ (Cmd⁷⁻⁵ = 000)

Logic₀ Logic₂ Cmd⁷⁻⁰ = RTS

When updating the PC, we must decide if the PC is to be replaced by a completely new value (as in the case of the JMP, JSR, JPI, JSI, JPZ and JSZ instructions), or by the result of the adder (all other instructions). The output from Logic₁ ANDed with Cycle can therefore be safely used by the multiplexor to obtain the new PC value (we need to always select PC+1 when Cycle is 0, even though we don't always write it to the PCA).

Note that the JPZ and JSZ instructions are implemented as 12 AND gates that cause the Accumulator value to be ignored, and the new PC to be set to 0. Likewise, the PCRamSel bit is cleared via these two instructions using the same AND mechanism.

The input to the 12-bit adder depends on whether we are incrementing by 1 (the usual case), or adding the offset as read from the MIU (when a branch is taken by the DBR and TBR instructions). Logic₃ generates the test.

Logic₃ Cycle

(((Cmd⁷⁻⁴ = DBR)

CZ)

((Cmd⁷⁻⁴ = TBR)

(Cmd₀ ⊕ Z)))

The actual offset to be added in the case of the DBR and TBR instructions is either the 8-bit value read from the MIU, or an 8-bit value generated by bits 3-1 of the opcode and treating bit 4 of the opcode as the sign (thereby making DBR immediate branching negative, and TBR immediate branching positive). The former is selected when bits 3-1 of the opcode is 0, as shown by Logic₄.

Logic₄ If (Cmd³⁻¹ = 000) output MIU8Data Else output Cmd₄|Cmd₄|Cmd₄|Cmd₄|Cmd₄|Cmd³⁻¹

Finally, the selection of which PC entry to use depends on the current value for SP. As we enter a subroutine, the SP index value must increment, and as we return from a subroutine, the SP index value must decrement. Logic₁ tells us when a subroutine is being entered, and Logic₂ tells us when the subroutine is being returned from. We use Logic₂ to select the altered SP value, but only write to the SP register when Exec and Cmd₄ are also set (to prevent JMP and JPZ from adjusting SP).

The two VAL units are validation units connected to the Tamper Prevention and Detection circuitry (described in Section 10.3.5 on page 962), each with an OK bit. The OK bit is set to 1 on PORstL, and ORed with the ChipOK values from both Tamper Detection Lines each cycle. The OK bit is ANDed with each data bit that passes through the unit. Both VAL units also parity-check the data bits to ensure that they are valid. If the parity-check fails, the Erase Tamper Detection Line is triggered.

In the case of VAL₁, the effective output from the SP register will always be 0. If the chip has been tampered with. This prevents an attacker from executing any subroutines.

In the case of VAL₂, the effective PC output will always be 0 if the chip has been tampered with. This prevents an attacker from executing any program code.

16 Address Generator Unit

The Address Generator Unit (AGU) generates effective addresses for accessing the Memory Unit (MU). In Cycle 0, the PC is passed through to the MU in order to fetch the next opcode. The AGU interprets the returned opcode in order to generate the effective address for Cycle 1. In Cycle 1, the generated address is passed to the MU.

The logic and registers contained in the AGU must be covered by both Tamper Detection Lines. This is to ensure that an attacker cannot alter any generated address. The latches for the counters and calculated address should also be parity-checked.

If either of the Tamper Detection Lines is broken, the AGU will generate address 0 each cycle and all counters will be fixed at 0. This will only come into effect if an attacker has disabled the RESET and/or erase circuitry, since under normal circumstances, breaking a Tamper Detection Line will result in a RESET or the erasure of all Flash memory.

16.1 Implementation

The block diagram for the AGU is shown in FIG. 404:

The accessMode and WriteMask registers must be cleared to 0 on reset to ensure that no access to memory occurs at startup of the CPU.

The Adr and accessMode registers are written to during the final cycle of cycle 0 (Fetch) and cycle 1 (Exec) with the address to use during the following cycle phase. For example, when cycle=1, the PC is selected so that it can be written to Adr during Exec. During cycle 0, while the PC is being output from Adr, the address to be used in the following cycle 1 is calculated (based on the fetched opcode seen as Cmd) and finally stored in Adr when Fetch is 1. The accessMode register is also updated in the same way.

It is important to distinguish between the value of Cmd during different values for Cycle:

-   -   During Cycle 0, when Fetch is 1, the 8-bit input Cmd holds the         instruction to be executed in the following Cycle 1. This 8-bit         value is used to decode the effective address for the operand of         the instruction.     -   During Cycle 1, when Exec is 1, Cmd holds the currently         executing instruction.

The WriteMask register is only ever written to during execution of an appropriate ROR instruction. Logic₁ sets the WriteMask and MMR WriteEnables respectively based on this condition:

Logic₁ Exec

(Cmd⁷⁻⁰ = ROR WriteMask)

The data written to the WriteMask register is the lower 8 bits of the Accumulator.

The Address Register Unit is only updated by an RIA or LIA instruction, so the writeEnable is generated by Logic₂ as follows:

Logic₂ Exec

(Cmd⁶⁻³ = 1111)

The Counter Unit (CU) generates counters C1, C2 and the selected N index. In addition, the CU outputs a CZ flag for use by the PCU. The CU is described in more detail below.

The VAL₁ unit is a validation unit connected to the Tamper Prevention and Detection circuitry (described in Section 10.3.5 on page 962). It contains an OK bit that is set to 1 on PORstL, and ORed with the ChipOK values from both Tamper Detection Lines each cycle. The OK bit is ANDed with the 12 bits of Adr before they can be used. If the chip has been tampered with, the address output will be always 0, thereby preventing an attacker from accessing other parts of memory. The VAL₁ unit also performs a parity check on the Adr Address bits to ensure it has not been tampered with. If the parity-check fails, the Erase Tamper Detection Line is triggered.

16.1.1 Counter Unit

The Counter Unit (CU) generates counters C1 and C2 (used internally). In addition, the CU outputs Cn and flag CZ for use externally. The block diagram for the CU is shown in FIG. 405: Registers C1 and C2 are updated when they are the targets of a DBR, SC or ROR instruction. Logic₁ generates the control signals for the write enables as shown in the following pseudocode.

isDbrSc = (Cmd⁷⁻⁴ = DBR)

(Cmd⁷⁻⁴ = SC) isRorCn = (Cmd⁷⁻⁴ = ROR)

(Cmd³⁻² = 10) CnWE = Exec

(isDbrSc

isRorCn) C1we = CnWE

Cmd₀ C2we = CnWE

Cmd₀

The single bit flag CZ is produced by the NOR of the appropriate C1 or C2 register for use during a DBR instruction. Thus CZ is 1 if the appropriate Cn value=0.

The actual value written to C1 or C2 depends on whether the ROR, DBR or SC instruction is being executed. During a DBR instruction, the value of either C1 or C2 is decremented by 1 (with wrap). One multiplexor selects between the lower 6 bits of the Accumulator (for ROR instructions), and a 6-bit value for an SC instruction where the upper 3 bits=the low 3 bits from C2, and low 3 bits=low 3 bits from Cmd. Note that only the lowest 3 bits of the operand are written to C1.

The two VAL units are validation units connected to the Tamper Prevention and Detection circuitry (described in Section 10.3.5 on page 962), each with an OK bit. The OK bit is set to 1 on PORstL, and ORed with the ChipOK values from both Tamper Detection Lines each cycle. The OK bit is ANDed with each data bit that passes through the unit. All VAL units also parity check the data to ensure the counters have not been tampered with. If a parity check fails, the Erase Tamper Detection Line is triggered.

In the case of VAL₁, the effective output from the counter C1 will always be 0 if the chip has been tampered with. This prevents an attacker from executing any looping constructs.

In the case of VAL₂, the effective output from the counter C2 will always be 0 if the chip has been tampered with. This prevents an attacker from executing any looping constructs.

16.1.2 Calculate Next Address

This unit generates the address of the operand for the next instruction to be executed. It makes use of the Address Register Unit and PC to obtain base addresses, and the counters from the Counter Unit to assist in generating offsets from the base address.

This unit consists of some simple combinatorial logic, including an adder that adds a 6-bit number to a 10-bit number. The logic is shown in the following pseudocode.

isErase = (Cmd⁷⁻⁰ = ERA) isSt = (Cmd⁷⁻⁴ = ST) isAccRead = (Cmd⁷⁻⁶ = 10) # First determine whether this is an immediate mode requiring PC+1 isJmpJsrDbrTbrImmed = (Cmd⁷⁻⁶ = 00)

(

Cmd₅

(Cmd⁵⁻¹ = 1x000)) isLia = (Cmd⁷⁻³ = LIA) isLogImmed = ((Cmd⁷⁻⁶ = 11)

((Cmd₅

Cmd₄)

(Cmd⁵⁻³ ≠ 111)))

(Cmd¹⁻⁰ = 10) pcSel = Cycle

(

Cycle

(isJmpJsrDbrTbrImmed

isLogImmed

isLia)) # Generate AnSel signal for the Address Register Unit A0Sel = (isAccRead

isSt)

(

Cmd₃

(Cmd⁵⁻³ = 001)) AnSel¹⁻⁰ =

A0Sel

Cmd²⁻¹ # The next address is either the new PC or must be generated # (we require the base address from Address Register Unit) nextRAMSel = AnDataOut₈

isErase If (nextRAMSel) baseAdr = 00 | AnDataOut⁷⁻⁰ # ram addresses are already word aligned Else baseAdr = AnDataOut⁷⁻⁰ | 00 # flash addresses are 4-byte aligned EndIf # Base address is now word (4-byte) aligned # Now generate the offset amount to be added to the base address selCn = (isAccRead

isSt)

(Cmd₅

Cmd₄)

Cmd₃ offset₀ = (A0Sel

Cmd₀)

(selCn

Cn₀) offset₁ = (A0Sel

Cmd₁)

(selCn

Cn₁) offset₂ = (A0Sel

Cmd₂)

(selCn

Cn₂) offset⁵⁻³ = selCn

Cn⁵⁻³ If (isErase) nextEffAdr¹¹⁻⁴ = Acc⁷⁻⁰ nextEffAdr³⁻⁰ = don't care Else # now we can simply add the offset to the base address to get the effective adr nextEffAdr¹¹⁻² = baseAdr + offset # 10 bit plus 6 bit, with wrap = 10 bits out nextEffAdr¹⁻⁰ = 0 # word access, so lower bits of effadr are 0 EndIf # Now generate the various signals for use during Cycle=1 # Note that these are only valid when pcSel is 0 (otherwise will read PC) nextAccessMode₀ = 1 # want 32-bit access nextAccessMode₁ = nextRAMSel # ram or flash access (only valid if rd/wr/erase set) nextAccessMode₂ = isAccRead # pcSel takes care of LIA instruction nextAccessMode₃ = isSt # write access nextAccessMode₄ = isErase # erase page access 16.1.3 Address Register Unit

This unit contains 4×9-bit registers that are optionally cleared to 0 on PORstL. The 2-bit input AnSel selects which of the 4 registers to output on DataOut. When the writeEnable is set, the AnSel selects which of the 4 registers is written to with the 9-bit DataIn.

17 Program Mode Unit

The Program Mode Unit (PMU) is responsible for Program Mode and Trim Mode operations:

-   -   Program Mode involves erasing the existing flash memory and         loading the new program/data into the flash. The program that is         loaded can be a bootstrap program if desired, and may contain         additional program code to produce a digital signature of the         final program to verify that the program was written correctly         (e.g. by producing a SHA-1 signature of the entire flash         memory).     -   Trim Mode involves counting the number of internal cycles that         have elapsed between the entry of Trim Mode (at the falling edge         of the ack) and the receipt of the next byte (at the falling         edge of the last bit before the ack) from the Master. When the         byte is received, the current count value divided by 2 is         transmitted to the Master.

The PMU relies on a fuse (implemented as the value of word 0 of the flash information block) to determine whether it is allowed to perform Program Mode operations. The purpose of this fuse is to prevent easy (or accidental) reprogramming of QA Chips once their purpose has been set. For example, an attacker may want to reuse chips from old consumables. If an attacker somehow bypasses the fuse check, the PMU will still erase all of flash before storing the desired program. Even if the attacker somehow disconnects the erasure logic, they will be unable to store a program in the flash due to the shadow nybbles.

The PMU contains an 8-bit buff register that is used to hold the byte being written to flash and a 12-bit adr register that is used to hold the byte address currently being written to.

The PMU is also used to load word 1 of the information block into a 32-bit register (combined from 8-bits of buff, 12-bits of adr, and a further 12-bit register) so it can be used to XOR all data to and from memory (both Flash and RAM) for future CPU accesses. This logic is activated only when the chip enters ActiveMode (so as not to access flash and possibly cause an erasure directly after manufacture since shadows will not be correct). The logic and 32-bit mask register is in the PMU to minimize chip area.

The PMU therefore has an asymmetric access to flash memory:

-   -   writes are to main memory     -   reads are from information block memory

The reads and writes are automatically directed appropriately in the MRU.

A block diagram of the PMU is shown in FIG. 406.

17.1 Local Storage and Counters

The PMU keeps a 1-cycle delayed version of MRURdy, called prevMRURdy. It is used to generate PMNewTrans. Therefore each cycle the PMU performs the following task:

-   -   prevMRURdy←MRURdy v(state=loadByte)

The PMU also requires 1-bit maskLoaded, idlePending and idlePending registers, all of which are cleared to 0 on RstL. The 1-bit fuseBlown register is set to 1 on RstL for security.

17.2 State Machine

The state machine for the PMU is shown in FIG. 407, with the pseudocode for the various states outlined below.

rstl

prevMRURdy, maskLoaded, idlePending, adr←0#clear most regs fuseBlown←1# for security sake assume the worst state←idle

The idle state, entered after reset, simply waits for the IOMode to enter ProgramMode, ActiveMode, or TrimMode. Note that the reset value for fuseBlown means that ProgramMode and TrimMode cannot be entered until after a successful entry into ActiveMode that also clears the fuseBlown register. In state idle, PMEn=

maskLoaded, and in state wait4Mode PMEn=0. In all other states, PMEn=1.

idle idlePending

0 PMEn =

maskLoaded PMNewTrans = 0 If ((IOMode = ActiveMode)

MRURdy) If (maskLoaded) state

wait4mode # no need to reload mask once loaded Else adr

0 # the location of the fuse is within 32-bit word 0 state

loadFuse EndIf ElseIf ((IOMode = ProgramMode)

MRURdy

fuseBlown) # wait 4 access 2 finish maskLoaded

0 # the mask is now invalid adr

0  # the location of the fuse is within 32-bit word 0 state

loadFuse ElseIf ((IOMode = TrimMode)

MRURdy

fuseBlown) # wait 4 access 2 finish maskLoaded

0 # the mask is now invalid adr

0 # start the counter on entering TrimMode state

trim Else state

idle EndIf

The wait4mode state simply waits until for the current mode to finish and returns to idle.

wait4mode PMEn = 0 PMNewTrans = 0 If (IOMode = IdleMode) state

idle Else state

wait4mode EndIf

The trim state is where we count the number of cycles between the entry of the Trim Mode and the arrival of a byte from the Master. When the byte arrives from the Master, we send the resultant count:

trim # We saturate the adder at all 1s to make external trim control easier lastOne = adr₀

adr₁

... adr₁₁ If (

lastOne) adr = adr + 1 # 12 bit incrementor EndIf # This logic simply causes the current adder value to be written to the # outByte when the inByte is received. The inByte is cleared when received # although it is not strictly necessary to do so PMOutByteWE = InByteValid # 0 in all other states PMInByteUsed = InByteValid # same as in loadByte state, 0 in all other states If (IOMode ≠ TrimMode) state

idle ElseIf (InByteValid) state

wait4mode Else state

trim EndIf

The loadFuse state is called whenever there is an attempt to program the device or we are entering ActiveMode and the mask is invalid (i.e. after power up or after a ProgramMode or TrimMode command). We load the 32-bit fuse value from word 0 of information memory in flash and compare it against the FuseSig constant (0x5555AAAA) to obtain the fuse value. The next state depends on IOMode and the Fuse.

loadFuse PMEn = 1 PMNewTrans = prevMRURdy idlePending_in = idlePending

(IOMode = IdleMode) idlePending

idlePending_in If (MRURdy) If (idlePending_in) # don't change state until the memory access is complete state

idle Else fuseBlown_in = (MRUData³¹⁻⁰ = FuseSig) fuseBlown

fuseBlown_in If (IOMode = ProgramMode) If (fuseBlown_in) state

wait4mode # not allowed to program anymore Else state

erase EndIf Elsif (IOMode = ActiveMode) adr

4 # byte 4 is word 1 (the location of the XORMask) state

getMask Else state

idle EndIf EndIf Else state

loadFuse EndIf

The erase state erases the flash memory and then leads into the main programming states:

erase PMNewTrans = prevMRURdy PMEraseDevice = 1 # is 0 in all other states adr

0 idlePending_in = idlePending

(IOMode ≠ ProgramMode) idlePending

idlePending_in If (MRURdy) If (idlePending_in) state

idle Else state

loadByte EndIf Else state

erase EndIf

Program Mode involves loading a series of 8-bit data values into the Flash. The PMU reads bytes via the IOU's InByte and InByteValID, setting MUInByteUsed as it loads data. The Master must send data slightly slower than the speed it takes to write to Flash to ensure that data is not lost.

loadByte # Load in 1 byte (1 word) from IO Unit PMNewTrans = 0 PMInByteUsed = InByteValid # same as in TrimIn state, and 0 in all other states If (IOMode ≠ ProgramMode) state

idle Else If (InByteValid) buff

InByte state

writeByte Else state

loadByte EndIf EndIf writeByte PMNewTrans = prevMRURdy PMRW = 0 # write. In all other states, PMRW = 1 (read) PM320ut⁷⁻⁰ = buff # data (can be tied to this) PM320ut¹⁹⁻⁸ = adr # can be tied to this PM320ut³¹⁻²⁰ = 12bitReg# is always this (is don't care during a write) idlePending_in = idlePending

(IOMode ≠ ProgramMode) idlePending

idlePending_in If (MRURdy) lastOne = adr₀

adr₁

... adr₁₁ adr

adr + 1 # 12 bit incrementor If (idlePending_in) state

idle ElseIf (lastOne) state

wait4Mode Else state

loadByte EndIf Else state

writeByte EndIf

The getMask state loads up word 1 of the flash information block (bytes 4-7) into the 32-bit buffer so it can be used to XOR all data to and from memory (both Flash and RAM) for future CPU accesses.

getMask PMNewTrans = prevMRURdy PM32Out¹⁹⁻⁸ = adr # adr should = 4, i.e. word 1 which holds the CPU's mask PMRW = 1 # read (MUST be 1 in this state) idlePending_in = idlePending

(IOMode ≠ ActiveMode) idlePending

idlePending_in If (MRURdy) buff

MRUData⁷⁻⁰ adr

MRUData¹⁹⁻⁸ 12bitReg

MRUData³¹⁻²⁰ maskLoaded

1 If (idlePending_in) state

idle Else state

wait4mode EndIf Else state

getMask EndIf 18 Memory Request Unit

The Memory Request Unit (MRU) provides arbitration between PMU memory requests and CPU-based memory requests.

The arbitration is straightforward: if the input PMEn is asserted, then PMU inputs are processed and CPU inputs are ignored. If PMEn is deasserted, the reverse is true.

A block diagram of the MRU is shown in FIG. 408.

18.1 Arbitration Logic

The arbitration logic block provides arbitration between the accesses of the PM and the 8/32-bit accesses of the CPU via a simple multiplexing mechanism based on PMEn:

ReqDataOut³¹⁻⁸ = CPUDataOut³¹⁻⁸ If (PMEn) NewTrans = PMNewTrans AccessMode₀ = PMRW # maps to 1 for reads (32 bits), 0 for writes (8 bits) AccessMode₁ = 0 # flash accesses only AccessMode₂ = PMRW

PMEraseDevice # read has lower priority than erase AccessMode₃ =

PMRW

PMEraseDevice # write has lower priority than erase AccessMode₄ = 0 # pageErase AccessMode₅ = PMEraseDevice # erase everything (main & info block) WriteMask = 0xFF Adr = PM32Out¹⁹⁻⁸ ReqDataOut⁷⁻⁰ = PM32Out⁷⁻⁰ Else NewTrans = CPUNewTrans

(CPUAccessMode⁴⁻² ≠ 000) AccessMode⁴⁻⁰ = CPUAccessMode AccessMode₅ = 0 # cpu cannot ever erase entire chip WriteMask = CPUWriteMask Adr = CPUAdr ReqDataOut⁷⁻⁰ = CPUDataOut⁷⁻⁰ EndIf 18.2 Memory Request Logic

The Memory Request Logic in the MRU implements the memory requests from the selected input.

An individual request may involve outputting multiple sub-requests e.g. an 8-bit read consists of 2×4-bit reads (each flash byte contains a nybble plus its inverse).

The input accessMode bits are interpreted as follows:

TABLE 382 Interpretation of accessMode bits Bit Description 0 0 = 8-bit access 1 = 32-bit access 1 0 = flash access 1 = RAM access this bit is only valid if bit 2, 3 or 4 is set 2 1 = read access 3 1 = write access 4 1 = erase page access 5 1 = erase entire (info and main) flash (only used within the MRU)

The MRU contains the following registers for general purpose flow control:

TABLE 383 Description of register settings name #bits Description ActiveTrans 1 Is there a transaction still running? If so, then extraTrans and nextToXfer can be considered valid. badUntilRestart 1 0 = memory (flash and ram) reads work correctly 1 = memory (flash and ram) reads return 0 Gets set whenever illChip gets set, and remains set until a soft restart occurs i.e. IOMode passes through Idle. extraTrans 1 Determines whether there is an additional sub- transaction to perform, e.g. a 32 bit read from flash involves 4 sub-transactions in the case of 8-bit accesses, and 8 sub-transactions in the case of 4-bit accesses. IllChip 1 0 = 15 consecutive bad reads have not occurred 1 = 15 consecutive bad reads have occurred nextToXfer 3 The next element (byte or nybble) number to transfer to/from memory restartPending 1 1 = IOMode passed through Idle while a transaction was being processed 0 = The transaction completed without IOMode passing through Idle retryCount 4 Number of times that a byte has been read badly from flash. When a byte has been read badly 15 consecutive times illChip will be set. retryStarted 1 0 = no retries encountered yet for this read 1 = retries have been encountered - retryCount holds the number of retries The retryStarted register is used to stop retryCount being cleared on good reads - thus keeping a record of the last number of retries on a bad read.

Table 383 lists the registers specifically for testing flash. Although the complete set of flash test registers is in both the MRU and MAU (group 0 is in the MRU, groups 1 and 2 are in the MAU), all the decoding takes place from the MRU.

TABLE 383 Flash test registers settable from CPU when the RAM address is >128⁷ adrbitSuperscriptparanumonly bits name description 0 0 shadowsOff 0 = regular shadowing (nybble based access to flash) 1 = shadowing disabled, 8-bit direct accesses to flash. 1 hiFlashAdr Only valid when shadowsOff = 1 0 = accesses are to lower 4 Kbytes of flash 1 = accesses are to upper 4 Kbytes of flash 2 1 3 enableFlashTest 0 = keep flash test register within the TSMC flash IP in its reset state 1 = enable flash test register to take on non-reset values.  8–4 flashTest Internal 5-bit flash test register within the TSMC flash IP (SFC008_08B9_HE). If this is written with 0x1E, then subsequent writes will be according to the TSMC write test mode. You must write a non-0x1E value or reset the register to exit this mode. 2 28–9 flashTime When timerSel is 1, this value is used for the duration of the program cycle within a standard flash write or erasure. 1 unit = 16 clock cycles (16 × 100 ns typical). Regardless of timerSel, this value is also used for the timeout following power down detection before the QA Chip resets itself. 1 unit = 1 clock cycle (=100 ns typical). Note that this means the programmer should set this to an appropriate value (e.g. 5 μs), just as the localId needs to be set. 29 timerSel 0 = use internal (default) timings for flash writes & erasures 1 = use flashTime for flash writes and erasures ⁷This is from the programmer's perspective. Addresses sent from the CPU are byte aligned, so the MRU needs to test bit n + 2. Similarly, checking DRAM address >128 means testing bit 7 of the address in the CPU, and bit 9 in the MRU. 18.2.1 Reset

Initialization on reset involves clearing all the flags:

MRURdy = 0  # can't process anything at this point activeTrans

0 extraTrans

0 illChip

0 badUntilRestart

0 restartPending

0 retryCount

0 retryStarted

0 nextToXfer

0  # don't care shadowsOff

0 hiFlashAdr

0 infoBlockSel

0# used to generate MRUMode₂ ⁸ unshadowed⁹ shadowed 18.2.2 Main Logic

The main logic consists of waiting for a new transaction, and starting an appropriate sub-transaction accordingly, as shown in the following pseudocode:

# Generate some basic signals for use in determining accessPatterns Is32Bit = AccessMode₀ Is8Bit =

AccessMode₀ IsFlash =

AccessMode₁ IsRAM = AccessMode₁ IsRead = AccessMode₂ IsWrite = AccessMode₃ noShadows = shadowsOff doShadows = IsFlash

noShadows continueRequest = (IOMode ≠ IdleMode) okForTrans =

restartPending

continueRequest startOfSubTrans = (NewTrans

extraTrans)

okForTrans doingTrans = startOfSubTrans

(activeTrans

extraTrans) IsInvalidRAM = doingTrans

IsRAM

(Adr₉

(Adr₈

Adr₇)) IsTestModeWE = doingTrans

IsRAM

IsWrite

Adr₉ IsTestReg₀ = IsTestModeWE

Adr₃ #write to flash test register - bit 1 of word adr IsTestReg₁ = IsTestModeWE

Adr₄ #write to flash test register - bit 2 of word adr MRUTestWE = IsTestReg₀

IsTestReg₁ IsPageErase = AccessMode₄ IsDeviceErase = AccessMode₅

(IsTestModeWE

(Adr⁸⁻² = 0001000)) # bit 9 not req IsErase = IsDeviceErase

IsPageErase MRURAMSel = IsRAM

MRUTestWE

IsDeviceErase IsInfBlock = (PMEn

(IsDeviceErase

IsRead))

(

PMEn

infoBlockSel

(IsDeviceErase

(IsFlash

(Adr¹¹⁻⁷ = 0)

(Adr₆

doShadows)))) # Which element (byte or nybble) are we up to xferring? If (NewTrans) toXfer = 0 Else toXfer = nextToXfer EndIf # Form the address that goes to the outside world If (IsFlash

noShadows) byteCount = toXfer¹⁻⁰ MRUAdr₁₂ = hiFlashAdr # upper or lower block of 4Kbytes of flash MRUAdr¹¹⁻² = Adr¹¹⁻² # word # MRUAdr¹⁻⁰ = (Adr¹⁻⁰

(

Is32Bit|

Is32Bit))

byteCount # byte Else byteCount = toXfer²⁻¹ MRUAdr¹²⁻³ = Adr¹¹⁻² # word # MRUAdr²⁻¹ = (Adr¹⁻⁰

(

Is32Bit|

Is32Bit))

byteCount # byte MRUAdr₀ = toXfer₀ #nybble EndIf # Assuming a write, are we allowed to write to this address? writeEn = SelectBit[WriteMask, ((MRUAdr₂

doShadows)| MRUAdr¹⁻⁰)]# mux: 1 from 8 # Generate the 4-bit mask to be used for XORing during CPU access to flash baseMask = SelectNybble(PM32Out, MRUAdr²⁻⁰) # mux selects 4 bits of 32 If (PMEn) theMask = 0 Else theMask = baseMask # we only use mask for CPU accesses to flash EndIf # Select a byte (and nybble) from the data for writes baseByte = SelectByte[ReqDataOut, byteCount] # mux: 8 bits from 32 baseNybble = SelectNybble[baseByte, toXfer₀] # mux: 4 bits from 8 outNybble = baseNybble ⊕ theMask # only used when nybble writing # Generate the data on the output lines (doesn't matter for reads or erasures) MRUDataOut³¹⁻⁸ = ReqDataOut³¹⁻⁸ # effectively don't care for flash writes If (doShadows) MRUDataOut₇ =

outNybble₃ MRUDataOut₆ = outNybble₃ MRUDataOut₅ =

outNybble₂ MRUDataOut₄ = outNybble₂ MRUDataOut₃ =

outNybble₁ MRUDataOut₂ = outNybble₁ MRUDataOut₁ =

outNybble₀ MRUDataOut₀ = outNybble₀ Else MRUDataOut⁷⁻⁰ = baseByte EndIf # Setup MRUMode allowTrans = IsRAM

IsRead

(IsWrite

writeEn)

IsErase If (doingTrans) MRUMode₂ = IsInfBlock MRUMode₁ = IsErase

IsTestReg₁ MRUMode₀ = IsDeviceErase 

 (

IsWrite 

 

IsPageErase) 

IsTestReg₀ MRUNewTrans = startOfSubTrans

allowTrans

(

IsInvalidRAM

MRUTestWE

IsDeviceErase) Else MRUMode²⁻⁰ = 001 # read (safe) MRUNewTrans = 0 EndIf # Generate the effective nybble read from flash (this may not be used). # When there is a shadowFault (non-erased memory and invalid shadows) we consider # it a bad read when an 8-bit read, or when writeMask₀ is 0. # Note: we always substitute the upper nybble of WriteMask for the non-valid data, # but only flag a read error if WriteMask₀ is also 1. When the data is erased, # we return 0 regardless of WriteMask₀. finishedTrans = doingTrans

MAURdy finishedFlashSubTrans = finishedTrans

IsFlash

IsErase isWrittenFlash = (FlashData⁷⁻⁰ ≠ 11111111) # flash is erased to all 1s If (isWrittenFlash

((FlashData_(7,5,3,1) ⊕ FlashData_(6,4,2,0)) ≠ 1111)) inNybble³⁻⁰ = WriteMask⁷⁻⁴ badRead = finishedFlashSubTrans

IsRead

(Is8Bit

WriteMask₀)

doShadows Else inNybble_(3,2,1,0) = (theMask_(3,2,1,0) ⊕ FlashData_(6,4,2,0))

isWrittenFlash badRead = 0 EndIf # Present the resultant data to the outside world MaskTheData = IsInvalidRAM

badRead

(badUntilRestart

IsRAM) NoData = IsErase

IsWrite

doingTrans If (NoData

MaskTheData) MRUData₀ = IsInvalidRAM

illChip MRUData⁴⁻¹ = retryCount

(IsInvalidRAM

Adr₂) # mask all 4 count bits MRUData³¹⁻⁵ = 0 # also ensures a read that is bad returns 0 ElseIf (IsRAM) MRUData³¹⁻²⁴ = SelectByte[RAMData, (Adr¹⁻⁰

Is32Bit|Is32Bit)] # mux: 8 from 32 MRUData²³⁻⁰ = RAMData²³⁻⁰ # lsbs remain unchanged from RAM ElseIf (doShadows) MRUData³¹⁻²⁸ = inNybble MRUData²⁷⁻⁰ = buff²⁷⁻⁰ Else MRUData³¹⁻²⁴ = FlashData MRUData²³⁻⁰ = buff²⁷⁻⁴ EndIf # Shift in the data for the good reads - either 4 or 8 bits (writes = don't care) If (finishedFlashSubTrans

badRead) buff³⁻⁰

buff⁷⁻⁴ # shift right 4 bits If (doShadows) buff²³⁻⁴

buff²⁷⁻⁸ # shift right 4 bits buff²⁷⁻²⁴

inNybble Else buff¹⁹⁻⁴

buff²⁷⁻¹² # shift right 8 bits, buff³⁻⁰ is don't care buff²⁷⁻²⁰

FlashData EndIf EndIf # Determine whether or not we need a new sub-transaction. We only need one if: # * there hasn't been a transition to IdleMode during this transaction # * we're doing 8 bit reads that are shadowed # * we're doing 32 bit reads and we've done less than 4 or 8 (sh vs non-sh) # * we got a bad read from flash and we need to retry the read (jic was a glitch) moreAdrsToGo = (

toXfer₀

((Is8Bit

doShadows)

Is32Bit))

(

toXfer₁

Is32Bit)

(

toXfer₂

Is32Bit

doShadows) needToRetryRead = badRead

(

retryStarted

(retryCount ≠ 1111)) extraTrans_in = finishedFlashSubTrans 

 (moreAdrsToGo 

needToRetryRead)

okForTrans nextToXfer 

 toXfer + (finishedFlashSubTrans 

 (IsWrite 

needToRetryRead)) # generate our rdy signal and state values for next cycle MRURdy =

doingTrans

(doingTrans

MAURdy

extraTrans_in) extraTrans

extraTrans_in activeTrans

MRURdy # all complete only when MRURdy is set # Take account of bad reads triedEnough = badRead

retryStarted

(retryCount = 1111) If (MAURdy) If (IsTestModeWE

(Adr⁵⁻² = 0000)) # capture writes to local regs illChip

0 retryCount

0 Else illChip

illChip

triedEnough If (badRead) retryCount

(retryCount

retryStarted) + 1 # AND all 4 bits retryStarted

1 Else retryStarted

0# clear flag so will be ok for the next read EndIf EndIf EndIf # Ensure that we won't have problems restarting a program If (MRURdy

okForTrans) # note MRURdy (may not be running a transaction!) shadowsOff, hiFlashAdr, infoBlockSel, restartPending, badUntilRestart

0 Else badUntilRestart

badUntilRestart

triedEnough If (doingTrans

ContinueRequest) restartPending

1 # record for later use EndIf If (IsTestModeWE

Adr₂) # the other writes are taken care of by the MAU shadowsOff

ReqDataOut₀ hiFlashAdr

ReqDataOut₁ infoBlockSel

ReqDataOut₂ EndIf EndIf 19 Memory Access Unit

The Memory Access Unit (MAU) takes memory access control signals and turns them into RAM accesses and flash access strobed signals with appropriate duration.

A new transaction is given by MRUNewTrans. The address to be read from or written to is on MRUAdr, which is a nybble-based address. The MRUAdr (13-bits) is used as-is for Flash addressing. When MRURAMSel=1, then the RAM address (RAMAdr) is taken from bits 9-3 of MRUAdr. The data to be written is on MRUData.

The return value MAURdy is set when the MAU is capable of receiving a new transaction the following cycle. Thus MAURdy will be 1 during the final cycle of a flash or ram access, and should be 1 when the MAU is idle. MAURdy should only be 0 during startup or when a transaction has yet to finish.

TABLE 384 Interpretation of MRUMode¹⁰ _(for RAM accesses) bits action xx0 doWrite xx1 doRead ¹⁰MRUMODE²⁻¹ is ignored for RAM accesses

When MRURAMSel=0, the access is to flash. If MRUTestWE=0, then the access is to regular flash memory, as given by MRUMode.

TABLE 385 Interpretation of MRUMode for regular flash accesses¹¹ bits 1-0 action when MRUMode₂ = 0 action when MRUMode₂ = 1 00 doWrite (main memory) doWrite (info block) 01 doRead (main memory) doRead (info block) 10 doErasePage (main memory) doErasePage (info block) 11 doEraseDevice (main memory) doEraseDevice (both blocks) ¹¹MRUMode₂ can be directly interpreted by the MAU as the IFREN signal required for embedded flash block SFC008_08B9_HE

If MRUTestWE is 1, then MRUMode₂ will also be 0, and the access is to a flash test register, as given by MRUMode:

TABLE 386 Interpretation of MRUMode for flash test register write accesses bits¹² action xx1 If (MRUData₃ = 0), tie the flash IP test register to its reset state If (MRUData₃ = 1), take the flash IP test register out of reset state, and write MRUData⁸⁻⁴ to the 5-bit flash test register within the flash IP (SFC008_08B9_HE) x1x Write MRUData²⁸⁻⁹ to the internal 20-bit alternate-counter- source register flashTime, and MRUData₂₉ to the corresponding 1-bit test register timerSel. ¹²MRUMode₂ will always be 0 when MRUTestWE = 1. 19.1 Implementation

The MAU consists of logic that calculates MAURdy, and additional logic that produces the various strobed signals according to the TSMC Flash memory SFC0008_(—)08B9_HE; refer to this datasheet [4] for detailed timing diagrams. Both main memory and information blocks can be accessed in the Flash. The Flash test modes are also supported as described in [5] and general application information is given in [6].

The MAU can be considered to be a RAM control block and a flash control block, with appropriate action selected by MRURAMSel. For all modes except read, the Flash requires wait states (which are implemented with a single counter) during which it is possible to access the RAM. Only 1 transaction may be pending while waiting for the wait states to expire. Multiple bytes may be written to Flash without exiting the write mode.

The MAU ensures that only valid control sequences meeting the timing requirements of the Flash memory are provided. A write time-out is included which ensures the Flash cannot be left in write mode indefinitely; this is used when the Flash is programmed via the IO Unit to ensure the X address does not change while in write mode. Otherwise, other units should ensure that when writing bytes to Flash, the X address does not change. The X address is held constant by the MAU during write and page erase modes to protect the Flash. If an X address change is detected by the MAU during a Flash write sequence, it will exit write mode allowing the X address to change and re-enter write mode. Thus, the data will still be written to Flash but it will take longer.

When either the Flash or RAM is not being used, the MAU sets the control signals to put the particular memory type into standby to minimise power consumption.

The MAU assumes no new transactions can start while one is in progress and all inputs must remain constant until MAU is ready.

19.2 Flash Test Mode

MAU also enables the Flash test mode register to be programmed which allows various production tests to be carried out. If MRUTestWE=1, transactions are directed towards the test mode register. Most of the tests use the same control sequences that are used for normal operation except that one time value needs to be changed. This is provided by the flashTime register that can be written to by the CPU allowing the timer to be set to a range of values up to more than 1 second. A special control sequence is generated when the test mode register is set to 0x1E and is initiated by writing to the Flash.

Note that on reset, timeSel and flashTime are both cleared to 0. The 5-bit flash test register within the TSMC flash IP is also reset by setting TMR=1. When MRUTestWE=1, any open write sequence is closed even if the write is not to the 5-bit flash test register within the TSMC flash IP.

19.3 Flash Power Failure Protection

Power could fail at any time; the most serious consequence would be if this occurred during writing to the Flash and data became corrupted in another location to that being written to. The MAU will protect the Flash by switching off the charge pump (high voltage supply used for programming and erasing) as soon as the power starts to fail. After a time delay of about 5 μs (programmable), to allow the discharge of the charge pump, the QA chip will be reset whether or not the power supply recovers.

19.4 Flash Access State Machine

19.5 Interface

TABLE 387 MAU interface description Signal name I/O Description Clk In System clock. RstL In System reset (active low). MAURAMEn In Flag indicating whether the external user needs access to the RAM at a gross level (e.g. the CPU is active and therefore may want RAM access). 1 = wants access available, 0 = don't want. MRUNewTrans In Flag indicating MRU wishes to start a new transaction. May only be asserted (=1) when MAURdy = 1. All inputs below must be held constant until MAU is ready. MRURAMSel In 1 = RAM, 0 = Flash. MRUMode2-0 In Type of transaction to be performed. MRUAdr12-0 In Memory address from the MRU. MRUDataOut31-0 In Data used to control and set test modes and timing. MRUTestWE In Flag indicating test mode transactions. PwrFailing In Flag indicating possible power failure in progress. MAURdy Out The MAU is ready when MAURdy = 1. It is always set for RAM transactions and held low during Flash wait states. RAMOutEn Out 0 = enable the RAM to read or write this cycle (i.e. active low) 1 = disable the RAM this cycle (saves power, memory is intact) RAMWE Out RAM write when RAMWE = 0 (Artisan Synchronous SRAM). MemClk Out Inverted system clock to the RAM (required to meet timing). FlashCtrl8-0 Out Control signals to the Flash. IFREN = information block enable, not used always = 0 XE = X address enable YE = Y address enable SE = sense amplifier enable (read only) OE = output enable (read only), hi-Z when OE = 0 PROG = program (write bytes) NVSTR = enables all write and erase modes ERASE = page erase mode MAS1 = mass erase mode TMR Out TMR = Register reset for test mode RAMAdr6-0 Out RAM address in the range 0 to 95. FlashAdr12-0 Out Flash address, full range. MAURstOutL Out Activates the global reset, RstL. 19.6 Calculation of Timer Values

Set and calculate timer initialisation values based on Flash data sheet values, clock period and clock range.

# Note: Flash data sheet gives minimum timings # Delays greater than 1 clock cycle clock_per = 100 # ns Flash_Tnvs = 7500 # ns Flash_Tnvh = 7500 # ns Flash_Tnvh1 = 150 # us Flash_Tpgs = 15 # us Flash_Tpgh = 100 # ns Flash_Tprog = 30 # us Flash_Tads = 100 # ns Flash_Tadh = 30 # us # Byte write timeout Flash_Trcv = 1500 # ns Flash_Thv = 6 # ms # Not currently used Flash_Terase = 30 # ms Flash_Tme = 300 # ms # Derive maximum counts (−1 since state machine is synchronous) FLASH_NVS = Flash_Tnvs/clock_per − 1 FLASH_NVH = Flash_Tnvh/clock_per − 1 FLASH_NVH1 = Flash_Tnvh1*1000/clock_per − 1 FLASH_PGS = Flash_Tpgs*1000/clock_per − 1 FLASH_PGH = Flash_Tpgh/clock_per − 1 FLASH_PROG = Flash_Tprog*1000/clock_per − 1 FLASH_ADS = Flash_Tads/clock_per − 1 FLASH_ADH = Flash_Tadh*1000/clock_per − 1 FLASH_ADH_AND_WRITE_PGH = FLASH_ADH + FLASH_PGH + 1 # note is + 1 FLASH_RCV = Flash_Trcv/clock_per − 1 FLASH_HV = Flash_Thv*1000000/clock_per − 1 FLASH_ERASE = Flash_Terase*1000000/clock_per − 1 FLASH_ME = Flash_Tme*1000000/clock_per − 1 count_size = 24 # Number of bits in timer counter (newCount) determined by Tme 19.7 Defaults

Defaults to use when no action is specified.

FlashTransPendingSet = 0 FlashTransPendingReset = 0 TMRSet = 0 TMRRst = 0 STLESet = 0 STLERst = 0 TestTimeEn = 0 IFREN = FlashXadr₇ XE = 0 YE = 0 SE = 0 OE = 0 PROG = 0 NVSTR = 0 ERASE = 0 MAS1 = 0 MAURstOutL = 1 If (accessCount ≠ 0) newCount =accessCount − 1 # decrement unless instructed otherwise Else newCount = 0 EndIf 19.8 Reset

Initialise state and counter registers

# asynchronous reset (active low) state

idle accessCount

1 countZ

0 XadrReg

0 FlashTransPending

0 TestTime

0 TMR

1 STLEFlag

0 19.9 State Machine

The state machine generates sequences of timed waveforms to control the operation of the Flash memory.

idle FlashTransPendingReset = 1 If (somethingToDo) # Flash starting conditions If (MRUTestWE) nextState =TM0 Else Switch (MRUModeint) Case doWrite: nextState =writeNVS newCount = FLASH_NVS Case doRead: YE = 1 SE = 1 OE = 1 XE = 1 nextState = idle Case doErasePage: nextState =pageErase newCount = FLASH_NVS Case doEraseDevice: nextState =massErase newCount = FLASH_NVS EndSwitch EndIf EndIf 19.9.1 Flash Page Erase

The following pseducocode illustrates the Flash page erase sequence.

pageErase ERASE = 1 XE = 1 If (

PwrFailing) If (countZ) newCount = FLASH_ERASE nextState =pageEraseERASE EndIf Else newCount = TestTime¹⁹⁻⁰ nextState =Help1 EndIf pageEraseERASE ERASE = 1 NVSTR = 1 XE = 1 If (

PwrFailing) If (countZ) newCount = FLASH_NVH nextState =pageEraseNVH EndIf Else newCount = TestTime¹⁹⁻⁰ nextState =Help1 EndIf pageEraseNVH NVSTR = 1 XE = 1 If (

PwrFailing) If (countZ) newCount = FLASH_RCV nextState =RCVPM EndIf Else newCount = TestTime¹⁹⁻⁰ nextState =Help1 EndIf RCVPM If (countZ) nextState =idle # exit EndIf 19.9.2 Flash Mass Erase

The following pseducocode illustrates the Flash mass erase sequence.

massErase MAS1 = 1 ERASE = 1 XE = 1 If (countZ) If (

TestTime₂₀) newCount = FLASH_ME Else newCount = TestTime¹⁹⁻⁰ | 0000 EndIf nextState =massEraseME EndIf massEraseME MAS1 = 1 ERASE = 1 NVSTR = 1 XE = 1 If (countZ) newCount = FLASH_NVH1 nextState =massEraseNVH1 EndIf massEraseNVH1 MAS1 = 1 NVSTR = 1 XE = 1 If (countZ) newCount = FLASH_RCV nextState = RCVPM EndIf 19.9.3 Flash Byte Write

The following pseducocode illustrates the Flash byte write sequence.

writeNVS PROG = 1 XE = 1 If (

PwrFailing) If (countZ) If (

STLEFlag) newCount = FLASH_PGS nextState =writePGS Else newCount = TestTime¹⁹⁻⁰ | 0000 nextState =STLE0 EndIf EndIf Else newCount = TestTime¹⁹⁻⁰ nextState =Help1 EndIf writePGS PROG = 1 NVSTR = 1 XE = 1 If (

PwrFailing) If (countZ) newCount = FLASH_ADS nextState =writeADS EndIf Else newCount = TestTime¹⁹⁻⁰ nextState =Help1 EndIf writeADS # Add Tads to Tpgs PROG = 1 NVSTR = 1 XE = 1 FlashTransPendingReset = 1 If (

PwrFailing) If (countZ) If (

TestTime₂₀) newCount = FLASH_PROG Else newCount = TestTime¹⁹⁻⁰ | 0000 EndIf nextState =writePROG EndIf Else newCount = TestTime¹⁹⁻⁰ nextState =Help1 EndIf writePROG PROG = 1 NVSTR = 1 YE = 1 XE = 1 If (

PwrFailing) If (countZ) newCount = FLASH_ADH_AND_WRITE_PGH nextState =writeADH EndIf Else newCount = TestTime¹⁹⁻⁰ nextState =Help2 EndIf writeADH PROG = 1 NVSTR = 1 XE = 1 FlashTransPendingSet = somethingToDo If (

PwrFailing) If (

FlashNewTrans) If (countZ) -- Gracefull exit after timeout newCount = FLASH_NVH nextState =writeNVH EndIf Else # -- Do something as there is a new transaction If ((MRUModeint = doWrite)

(

XadrCh)) newCount = FLASH_ADS -- Write another byte nextState =writeADS Else newCount = FLASH_NVH -- Exit as new trans is not Flash write nextState =writeNVH EndIf EndIf Else newCount = TestTime¹⁹⁻⁰ nextState =Help1 EndIf writeNVH NVSTR = 1 XE = 1 FlashTransPendingSet = somethingToDo If (

PwrFailing) If (countZ) newCount = FLASH_RCV nextState =RCV EndIf Else newCount = TestTime¹⁹⁻⁰ nextState =Help1 EndIf RCV # wait til we're allowed to do another transaction FlashTransPendingSet = somethingToDo If (countZ) nextState =idle EndIf 19.9.4 Test Mode Sequence

The following pseducocode illustrates the test mode sequence.

TM0 # Needed this due to delay on TMR IFREN = 0 nextState = idle # default If ( MRUModeint₁) TestTimeEn = 1 EndIf If (MRUModeint₀) If (

MRUDataOut₃) TMRSet = 1 STLERst = 1 # Reset flag as leaving test mode Else If (MRUDataOut⁸⁻⁴ = 11110) STLESet = 1 Else STLERst = 1 EndIf TMRRst = 1 nextState =TM1 # Will get priority EndIf EndIf TM1 IFREN = 0 nextState =TM2 TM2 NVSTR = 1 SE = 1 IFREN = 0 nextState =TM3 TM3 NVSTR = 1 SE = 1 MAS1 = MRUDataOut₄ IFREN = MRUDataOut₅ XE = MRUDataOut₆ YE = MRUDataOut₇ ERASE = MRUDataOut₈ TMRSet = 1 nextState =TM4 TM4 NVSTR = 1 SE = 1 MAS1 = MRUDataOut₄ IFREN = MRUDataOut₅ XE = MRUDataOut₆ YE = MRUDataOut₇ ERASE = MRUDataOut₈ TMRRst = 1 nextState =TM5 TM5 NVSTR = 1 SE = 1 MAS1 = MRUDataOut₄ IFREN = MRUDataOut₅ XE = MRUDataOut₆ YE = MRUDataOut₇ ERASE = MRUDataOut₈ nextState =TM6 TM6 NVSTR = 1 SE = 1 nextState =idle 19.9.5 Reverse Tunneling and Thin Oxide Leak Test

The following pseducocode shows the reverse tunneling and thin oxide leak test sequence.

STLE0 XE = 1 PROG = 1 NVSTR = 1 If (countZ) newCount = FLASH_NVH nextState =STLE1 EndIf STLE1 XE = 1 NVSTR = 1 If (countZ) newCount = FLASH_RCV nextState =STLE2 EndIf STLE2 If (countZ) nextState =idle EndIf 19.9.6 Emergency Instructions

The following pseducocode shows the states used for emergency situations such as when power is failing.

Help1 # MAURdy −> 0 to hold MAU inputs constant, if not too late XE = 1 If (countZ) nextState =Goodbye EndIf Help2 # MAURdy −> 0 to hold MAU inputs constant, if not too late XE = 1 YE = 1 If (countZ) nextState =Goodbye EndIf Goodbye XE = 1 # Prevents Flash timing violation MAURstOutL = 0 # Reset whole chip whether power fails # nothing else to do or recovers 19.10 Concurrent Logic

accessCount

newCount # update accessCount every cycle countZ

(newCount = 0) XadrReg

FlashXAdr # store the previous X address state

nextState If (FlashTransPendingReset) FlashTransPending

0 # Reset flag (has priority) Else If (FlashTransPendingSet) FlashTransPending

1 # Set flag EndIf EndIf If (TestTimeEn) TestTime

MRUDataOut²⁹⁻⁹ EndIf If (TMRSet) -- SRFF for TMR TMR

1 Else If (TMRRst) TMR

0 EndIf EndIf If (STLERst) -- SRFF for STLE tests STLEFlag

0 Else If (STLESet) STLEFlag

1 EndIf EndIf FlashNewTrans = MRUNewTrans

(

MRURAMSel) RAMNewTrans = MRUNewTrans

MRURAMSel somethingToDo = FlashTransPending

FlashNewTrans quickCmd = (MRUModeint = doRead)

MRUTestWE FlashRdy = ((state = idle)

(

somethingToDo

quickCmd))

(((state = writeADH)

(state = writeNVH)

(state = writeRCV))

(

FlashTransPendingSet))

((state = TM0)

(nextState = idle))

(state = TM6) If (MRURamSel) MAURdy = 1 # Always ready for RAM Else MAURdy = FlashRdy EndIf IandX = MRUMode₂ | MRUAdr¹²⁻⁶ FlashXAdr = IandX When ((

XE)

(SE

OE)) Else XadrReg FlashAdr = FlashXAdr | MRUAdr⁵⁻⁰ # Merge X and Y addresses XadrCh = 1 When ((XadrReg /= IandX)

 XE 

(

SE) 

(

OE)

FlashNewTrans) Else 0 # Xadr change MRUModeint = MRUMode¹⁻⁰ # Backwards compatability RAMAdr = MRUAdr⁹⁻³ # maximum address = 95, responsibility of MRU for valid adr RAMWE = MRUModeint₀ RAMOutEn =

RAMNewTrans # turn off RAM if not using it FlashCtrl(0) = IFREN FlashCtrl(1) = XE FlashCtrl(2) = YE FlashCtrl(3) = SE FlashCtrl(4) = OE FlashCtrl(5) = PROG FlashCtrl(6) = NVSTR FlashCtrl(7) = ERASE FlashCtrl(8) = MAS1 MemClk =

Clk # Memory clock 20 Analogue Unit

This section specifies the mandatory blocks of Section 11.1 on page 965 in a way which allows some freedom in the detailed implementation.

Circuits need to operate over the temperature range −40° C. to +125° C.

The unit provides power on reset, protection of the Flash memory against erroneous writes during power down (in conjunction with the MAU) and the system clock SysClk.

20.1 Voltage Budget

The table below shows the key thresholds for V_(DD) which define the requirements for power on reset and normal operation.

TABLE 388 V_(DD) limits VDD parameter Description Voltage VDDFTmax Flash test maximum 3.6¹³ VDDFTtyp Flash test typical 3.3 VDDFTmin Flash test minimum 3.0 VDDmax Normal operation maximum (typ + 10%) 2.75¹⁴ VDDtyp Normal operation typical 2.5 VDDmin Normal operation minimum (typ − 5%) 2.375 VDDPORmax Power on reset maximum 2.0¹⁵ ¹³The voltage VDDFT may only be applied for the times specified in the TSMC Flash memory test document. ¹⁴Voltage regulators used to derive VDD will typically have symmetric tolerance limits ¹⁵The minimum allowable voltage for Flash memory operation. 20.2 Voltage Reference

This circuit generates a stable voltage that is approximately independent of PVT (process, voltage, temperature) and will typically be implemented as a bandgap. Usually, a startup circuit is required to avoid the stable V_(bg)=0 condition. The design should aim to minimise the additional voltage above V_(bg) required for the circuit to operate. An additional output, BGOn, will be provided and asserted when the bandgap has started and indicates to other blocks that the output voltage is stable and may be used.

TABLE 389 Bandgap target performance Parameter Conditions Min Typ Max Units Vbg¹⁶ typical 1.2 1.23 1.26 V IDD typical 50 μA Vstart worst case 1.6 V Iout 10 nA Vtemp +0.1 mV/o C. ¹⁶Over PVT, not including offsets 20.3 Power Detection Unit

Only under voltage detection will be described and is required to provide two outputs:

-   -   underL controls the power on reset; and     -   PwrFailing indicates possible failure of the power supply.

Both signals are derived by comparing scaled versions of V_(DD) against the reference voltage V_(bg).

20.3.1 V_(DD) Monotonicity

The rising and falling edges of V_(DD) (from the external power supply) shall be monotonic in order to guarantee correct operation of power on reset and power failing detection. Random noise may be present but should have a peak to peak amplitude of less than the hysteresis of the comparators used for detection in the PDU.

20.3.2 Under Voltage Detection Unit

The underL signal generates the global reset to the logic which should be de-asserted when the supply voltage is high enough for the logic and analogue circuits to operate. Since the logic reset is asynchronous, it is not necessary to ensure the clock is active before releasing the reset or to include any delay.

The QA chip logic will start immediately the power on reset is released so this should only be done when the conditions of supply voltage and clock frequency are within limits for the correct operation of the logic.

The power on reset signal shall not be triggered by narrow spikes (<100 ns) on the power supply. Some immunity should be provided to power supply glitches although since the QA chip may be under attack, any reset delay should be kept short. The unit should not be triggered by logic dynamic current spikes resulting in short voltage spikes due to bond wire and package inductance. On the rising edge of V_(DD), the maximum threshold for de-asserting the signal shall be when V_(DD)>V_(DDmin). On the falling edge of V_(DD), the minimum threshold for asserting the signal shall be V_(DD)<V_(DDPORmax).

The reset signal must be held low long enough (T_(pwmin)) to ensure all flip-flops are reset. The standard cell data sheet [7] gives a figure of 0.73 ns for the minimum width of the reset pulse for all flip-flop types.

2 bits of trimming (trim₁₋₀) will be provided to take up all of the error in the bandgap voltage. This will only affect the assertion of the reset during power down since the power on default setting must be used during power up.

Although the reference voltage cannot be directly measured, it is compared against V_(DD) in the PDU. The state of the power on reset signal can be inferred by trying to communicate through the serial bus with the chip. By polling the chip and slowly increasing V_(DD), a point will be reached where the power on reset is released allowing the serial bus to operate; this voltage should be recorded. As V_(DD) is lowered, it will cross the threshold which asserts the reset signal. The power on default is set to the lowest voltage that can be trimmed (which gives the maximum hysterisis). This voltage should be recorded (or it may be sufficient to estimate it from the reset release voltage recorded above). V_(DD) is then increased above the reset release threshold and the PDU trim adjusted to the setting the closest to V_(DDPORmax). V_(DD) should then be lowered and the threshold at which the reset is re-asserted confirmed.

TABLE 390 Power on reset target performance Parameter Conditions Min Typ Max Units Vthrup T = 27o C. 2.0 2.375 V Vthrdn T = 27o C. 2.0 2.1 V Vhystmin 16 mV IDD 5 μA Tspike 100 ns Vminr 0.5 V Tpwmin 1 ns Power on Reset Behaviour

The signal PwrFailing will be used to protect the Flash memory by turning off the charge pump during a write or page erase if the supply voltage drops below a certain threshold. The charge pump is expected to take about 5 us to discharge. The PwrFailing signal shall be protected against narrow spikes (<100 ns) on the power supply.

The nominal threshold for asserting the signal needs to be in the range V_(PORmax)<V_(DDPFtyp)<V_(DDmin) so is chosen to be asserted when V_(DD)<V_(DDPFtyp)=V_(DDPORmax)+200 mV. This infers a V_(DD) slew rate limitation which must be <200 mV/5 us to ensure enough time to detect that power is failing before the supply drops too low and the reset is activated. This requirement must be met in the application by provision of adequate supply decoupling or other means to control the rate of descent of V_(DD).

TABLE 391 Power failing detection target performance Parameter Conditions Min Typ Max Units Vthr T = 27o C. 2.1 2.2 2.3 V¹⁷ Vhyst 16 mV IDD 5 μA Tspike 100 ns Vminr 0.5 V ¹⁷These limits are after trimming and include an allowance for VDD ramping.

2 bits of trimming (trim₁₋₀) will be provided to take up all of the error in the bandgap voltage.

20.4 Ring Oscillator

SysClk is required to be in the range 7-14 MHz throughout the lifetime of the circuit provided V_(DD) is maintained within the range V_(DDMIN)<V_(DD)<V_(DDMAX). The 2:1 range is derived from the programming time requirements of the TSMC Flash memory. If this range is exceeded, the useful lifetime of the Flash may be reduced.

The first version of the QA chip, without physical protection, does not require the addition of random jitter to the clock. However, it is recommended that the ring oscillator be designed in such a way as to allow for the addition of jitter later on with minimal modification. In this way, the un-trimmed centre frequency would not be expected to change.

The initial frequency error must be reduced to remain within the range 10 MHz/1.41 to 10 MHz×1.41 allowing for variation in:

-   -   voltage     -   temperature     -   ageing     -   added jitter     -   errors in frequency measurement and setting accuracy

The range budget must be partitioned between these variables.

FIG. 411._Ring oscillator block diagram

The above arrangement allows the oscillator centre frequency to be trimmed since the bias current of the ring oscillator is controlled by the DAC. SysClk is derived by dividing the oscillator frequency by 5 which makes the oscillator smaller and allows the duty cycle of the clock to be better controlled.

20.4.1 DAC (Programmable Current Source)

Using V_(bg), this block sources a current that can be programmed by the Trim signal. 6 of the available 8 trim bits will be used (trim₇₋₂) giving a clock adjustment resolution of about 250 kHz. The range of current should be such that the ring oscillator frequency can be adjusted over a 4 to 1 range.

TABLE 392 Programmable current source target performance Parameter Conditions Min Typ Max Units Iout Trim7 − 2 = 0  5 μA Trim7 − 2 = 32 12.5 Trim7 − 2 = 63 20 Vrefin 1.23 V Rout Trim7 − 2 = 63 2.5 MΩ 20.4.2 Ring Oscillator Circuit

TABLE 393 Ring oscillator target performance Parameter Conditions Min Typ Max Units Fosc¹⁸ 7 10 14 MHz IDD 10 μA KI 1 MHz/μA KVDD +200 KHz/V KT +30 KHz/o C. Vstart 1.5 V K_(I) = control sensitivity, K_(VDD) = V_(DD) sensitivity, K_(T) = temperature sensitivity With the figures above, K_(VDD) will give rise to a maximum variation of ±50 kHz and K_(T) to ±1.8 MHz over the specified range of V_(DD) and temperature. ¹⁸Accounting for division by 5 20.4.3 Div5

The ring oscillator will be prescaled by 5 to obtain the nominal 10 MHz clock. An asynchronous design may be used to save power. Several divided clock duty cycles are obtainable, eg 4:1, 3:2 etc. To ease timing requirements for the standard cell logic block, the following clock will be generated; most flip-flops will operate on the rising edge of the clock allowing negative edge clocking to meet memory timing.

TABLE 394 Div5 target performance Parameter Conditions Min Typ Max Units Fmax Vdd = 1.5 V 100 MHz IDD 10 μA 20.5 Power on Reset

This block combines the overL (omitted from the current version), underL and MAURstOutL signals to provide the global reset. MAURstOutL is delayed by one clock cycle to ensure a reset generated when this signal is asserted has at least this duration since the reset deasserts the signal itself. It should be noted that the register, with active low reset RN, is the only one in the QA chip not connected to RstL.

[4] TSMC, Oct. 1, 2000, SFC0008_(—)08B9_HE, 8K×8 Embedded Flash Memory Specification, Rev 0.1.

[5] TSMC (design service division), Sep. 10, 2001, 0.25 um Embedded Flash Test Mode User Guide, V0.3.

[6] TSMC (EmbFlash product marketing), Oct. 19, 2001, 0.25 um Application Note, V2.2.

[7] Artisan Components, January 1999, Process Perfect Library Databook 2.5-Volt Standard Cells, Rev1.0.

Other Applcations for Protocols and QA Chips

1 Introduction

In its preferred form, the QA chip [1] is a programmable 32 bit microprocessor with security features (8,000 gates, 3 k bits of RAM and 8 kbytes of flash memory for program and non-volatile data storage). It is manufactured in a 0.25 um CMOS process.

Physically, the chip is mounted in a 5 pin SOT23 plastic package and communicates with external circuitry via a two pin serial bus.

The QA chip was designed to for authenticating consumable usage and performance upgrades in printers and associated hardware.

Because of its core functionality and programmability the QA chip can also be used in applications that differ significantly from its original one. This document seeks to identify some of those areas.

3 Applications Overview

Applications include:

-   -   Regular EEPROM     -   Secure EEPROM     -   General purpose MPU with security features     -   Security coprocessor for microprocessor system     -   Security coprocessor for PC (with optional USB connection)     -   Resource dispenser—secure, web based transfer of a variable         quantity from “source” to “sink”     -   IDtag     -   Security pass inside offices     -   Set top box security     -   Car key     -   Car Petrol     -   Car manufacturer “genuine parts” detection, where the car         requires genuine (or authorised) parts to function.     -   Aeroplane control on motor-control servos to allow secure         external control on an aircraft in a hijack situation.     -   Security device for controlling access to and copying of audio,         video, and data (eg, preventing unauthorized downloading of         music to a device).         4 Exemplary Application Descriptions         4.1 Car Petrol

Using mechanisms and protocols similar to those described in relation to ink refills, refilling of petrol can be controlled. An example of a commercial relationship this allows is selling a car at a discounted rate, but requiring that the car be refilled at designated service stations. Similarly, prevention of unauthorized servicing can be achieved.

4.2 Car Keys

4.2.1 Basic Advantages Over Physical Keys

-   -   Keys and locks can be easily programmed & configured for use     -   Can only be duplicated/reprogrammed by an authorised individual     -   The same key can be used for physical entry/exit and remote         (radio-based) entry/exit     -   Inbuilt security features         4.2.2 Single Key for Multiple Vehicles

Useful when a family has more than one car.

-   -   Can be programmed so any keys fits any car.     -   Fewer number of duplicate keys.     -   Misplacing a key for a particular car—any key for any other car         can be used as oppose to duplicate of the same key.         4.2.3 Multiple Keys for a Single Vehicle         4.2.3.1 Same Company Car Being Driven by Multiple Drivers     -   Mileage can be logged per driver e.g. for accounting purposes.     -   Key permissions can be different per driver (e.g. boot/trunk         access may be disabled)         4.2.3.2 Same Family Car Being Driven by Children and Parents     -   Time/date restrictions can be applied to (e.g. children's) keys     -   Speeds above a specified limit (and duration of that speed) can         be logged for auditing purposes (may be less dangerous than         actually enforcing a speed limit)         4.2.4 No Problem if Key Lost

Can easily:

-   -   make a new key the same as lost one (existing copies of key will         still function)     -   reprogram the locks on car (and reprogram all non-lost keys to         match) so the lost key will no longer function         4.2.5 No Problem if Key Left in Car     -   Easy to create a one-time-use open-door-only key via roadside         assistance based on secret password information, driver's         license etc (prevents having to break into the car)         4.2.6 Car Rentals     -   Key can have an expiration date (e.g. some period past the         rental end-date)         4.2.7 Single Physical Key for All Locks in Car

A single physical key can open all locks (door, immobiliser, boot/trunk, glovebox etc.).

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1. A method for arbitrating between a plurality of access requests issued in relation to a resource by a plurality of requestors in an integrated circuit, wherein each request can be one of a CPU write request from a CPU requester, a non-CPU write request from a non-CPU requester and a non-CPU read request from a non-CPU requester, the method including the steps of: (a) receiving, in a timeslot arbitrator of the integrated circuit, a plurality of the access requests; (b) maintaining, in the timeslot arbitrator, a current pointer that points to a current timeslot in a timeslot list, and at least one lookahead pointer that points to a future timeslot in the timeslot list; and (c) in the event an access request as arbitrated via the lookahead pointer is a non-CPU write request, initiating performance of the access request, in the timeslot arbitrator, earlier than the position in the list suggests it would be performed should it be started when the current pointer reached the current timeslot, wherein, in step (c), the earlier position is selected so as to not be adjacent a position in the list for performance of another non-CPU write request, and each timeslot in the timeslot list is configured with a CPU access preceding a non-CPU access so that a CPU write request is performed before either a non-CPU write request or non-CPU read request.
 2. A method according to claim 1, wherein step (c) includes the substep of performing the access request indicated by the lookahead pointer immediately after the access request indicated by the current pointer is performed.
 3. A method according to claim 1, wherein step (c) includes the timeslot arbitrator arbitrating in the timeslot list CPU write requests to be interleaved with non-CPU write and read requests.
 4. A method according to claim 3, wherein step (c) includes the substep of performing the access request indicated by the lookahead pointer immediately after the access request indicated by the current pointer is performed.
 5. A method according to claim 1, wherein the number of timeslots between the timeslot indicated by the lookahead pointer and the timeslot indicated by the current pointer takes into account a latency difference between performing the access requests.
 6. An integrated circuit including: a plurality of operative units, each of which is capable of issuing a request for access to a memory accessible by the integrated circuit including a CPU write request from a CPU requester, a non-CPU write request from a non-CPU requester and a non-CPU read request from a non-CPU requester; and an timeslot arbitrator for arbitrating between requests issued by the operative units for access to the memory, the timeslot arbitrator being configured to: (a) receive a plurality of the access requests; (b) maintain a current pointer that points to a current timeslot in a timeslot list, and at least one lookahead pointer that points to a future timeslot in the timeslot list; and (c) in the event the access request as arbitrated via the lookahead pointer is a non-CPU write request, performing the access request earlier than the position in the list suggests it should be performed should it be started when the current pointer reached the current timeslot, wherein, in (c), the earlier position is selected so as to not be adjacent a position in the list for performance of another non-CPU write request, and each timeslot in the timeslot list is configured with a CPU access preceding a non-CPU access so that a CPU write request is performed before either a non-CPU write request or non-CPU read request.
 7. An integrated circuit according to claim 6, wherein in (c) the timeslot arbitrator arbitrates, in the timeslot list, CPU write requests to be interleaved with non-CPU write and read requests.
 8. An integrated circuit according to claim 6, wherein the number of timeslots between the timeslot indicated by the lookahead pointer and the timeslot indicated by the current pointer takes into account a latency difference between performing the access requests. 